[计算机网络系统方法].(美)peterson.清晰版


THIRD EDITION COMPUTER NETWORKS A Systems Approach The Morgan Kaufmann Series in Networking Series Editor, David Clark, M.I.T. Computer Networks: A Systems Approach, 3e Larry L. Peterson and Bruce S. Davie Network Architecture, Analysis, and Design, 2e James D. McCabe MPLS Network Management: MIBs, Tools, and Techniques Thomas D. Nadeau Developing IP-Based Services: Solutions for Service Providers and Vendors Monique Morrow and Kateel Vijayananda Telecommunications Law in the Internet Age Sharon K. Black Optical Networks: A Practical Perspective, 2e Rajiv Ramaswami and Kumar N. Sivarajan Internet QoS: Architectures and Mechanisms Zheng Wang TCP/IP Sockets in Java: Practical Guide for Programmers Michael J. Donahoo and Kenneth L. Calvert TCP/IP Sockets in C: Practical Guide for Programmers Kenneth L. Calvert and Michael J. Donahoo Multicast Communication: Protocols, Programming, and Applications Ralph Wittmann and Martina Zitterbart MPLS: Technology and Applications Bruce Davie and Yakov Rekhter High-Performance Communication Networks, 2e Jean Walrand and Pravin Varaiya Internetworking Multimedia Jon Crowcroft, Mark Handley, and Ian Wakeman Understanding Networked Applications: A First Course David G. Messerschmitt Integrated Management of Networked Systems: Concepts, Architectures, and their Operational Application Heinz-Gerd Hegering, Sebastian Abeck, and Bernhard Neumair Virtual Private Networks: Making the Right Connection Dennis Fowler Networked Applications: A Guide to the New Computing Infrastructure David G. Messerschmitt Modern Cable Television Technology: Video, Voice, and Data Communications Walter Ciciora, James Farmer, and David Large Switching in IP Networks: IP Switching, Tag Switching, and Related Technologies Bruce S. Davie, Paul Doolan, and Yakov Rekhter Wide Area Network Design: Concepts and Tools for Optimization Robert S. Cahn Frame Relay Applications: Business and Technology Case Studies James P. Cavanagh For further information on these books and for a list of forthcoming titles, please visit our website at http://www.mkp.com THIRD EDITION Larry L. Peterson & Bruce S. Davie COMPUTER NETWORKS A Systems Approach Senior Editor Rick Adams Publishing Services Manager Simon Crump Developmental Editor Karyn Johnson Cover Design Ross Carron Design Cover Image Vasco de Gama Bridge, Lisbon, Portugal Composition/Illustration International Typesetting and Composition Copyeditor Ken DellaPenta Proofreader Jennifer McClain Indexer Steve Rath Printer Courier Corporation Designations used by companies to distinguish their products are often claimed as trademarks or registered trademarks. In all instances in which Morgan Kaufmann Publishers is aware of a claim, the product names appear in initial capital or all capital letters. Readers, however, should contact the appropriate companies for more complete information regarding trademarks and registration. Morgan Kaufmann Publishers An Imprint of Elsevier Science 340 Pine Street, Sixth Floor San Francisco, CA 94104-3205 www.mkp.com © 2003 by Elsevier Science (USA) All rights reserved Printed in the United States of America 0706050403 54321 No part of this publication may be reproduced, stored in a retrieval system, or transmitted in any form or by any means—electronic, mechanical, photocopying, or otherwise—without the prior written permission of the publisher. Library of Congress Control Number: xxxxxxxxxx ISBN: 1-55860-832-X (Casebound) ISBN: 1-55860-833-8 (Paperback) This book is printed on acid-free paper. To Lee Peterson and Robert Davie This Page Intentionally Left Blank FOREWORD David Clark Massachusetts Institute of Technology This third edition represents another major upgrade to this classic networking book. The field continues to change fast, and new concepts emerge with amaz- ing speed. This version expands its discussion of a lot of important new top- ics, including peer-to-peer networks, Ipv6, overlay and content distribution networks, MPLS and switching, wireless and mobile technology, and more. It also contains an earlier and stronger focus on applications, which reflects the student and professional’s increased familiarity with a wide range of networked applications. The book continues its tradition of giving you the facts you need to understand today’s world. But it has not lost track of its larger goal, to tell you not only the facts but the why behind the facts. The philosophy of the book remains the same: to be timely but timeless. What this book will teach you in today’s networked world will give you the insight needed to work in tomorrow’s landscape. And that is important, since there is no reason to believe that the evolution of networks is going to slow down anytime soon. It is hard to remember what the world looked like only ten years ago. Back then the Internet was not really a commercial reality. Ten megabits per second was really fast. We didn’t worry about spam and virus attacks—we left our computers unguarded and hardly worried. Those times were simpler, but today may be more exciting. And you better believe that tomorrow will be different from today: at least as exciting, with luck no less trustworthy, and certainly bigger, faster and filled with fresh innovation. So I hope Larry and Bruce can relax for a little before they have to start the next revision. Meanwhile, use this book to learn about today and get ready for tomorrow. Have fun. This Page Intentionally Left Blank FOREWORD TO THE FIRST EDITION David Clark Massachusetts Institute of Technology The term spaghetti code is universally understood as an insult. All good computer scientists worship the god of modularity, since modularity brings many benefits, including the all-powerful benefit of not having to understand all parts of a problem at the same time in order to solve it. Modularity thus plays a role in presenting ideas in a book, as well as in writing code. If a book’s material is organized effectively— modularly—the reader can start at the beginning and actually make it to the end. The field of network protocols is perhaps unique in that the “proper” modularity has been handed down to us in the form of an international standard: the seven-layer reference model of network protocols from the ISO. This model, which reflects a layered approach to modularity, is almost universally used as a starting point for discussions of protocol organization, whether the design in question conforms to the model or deviates from it. It seems obvious to organize a networking book around this layered model. However, there is a peril to doing so, because the OSI model is not really successful at organizing the core concepts of networking. Such basic requirements as reliability, flow control, or security can be addressed at most, if not all, of the OSI layers. This fact has led to great confusion in trying to understand the reference model. At times it even requires a suspension of disbelief. Indeed, a book organized strictly according to a layered model has some of the attributes of spaghetti code. Which brings us to this book. Peterson and Davie follow the traditional layered model, but they do not pretend that this model actually helps in the understanding of the big issues in networking. Instead, the authors organize discussion of fundamental concepts in a way that is independent of layering. Thus, after reading the book, readers will understand flow control, congestion control, reliability enhancement, data rep- resentation, and synchronization, and will separately understand the implications of addressing these issues in one or another of the traditional layers. This is a timely book. It looks at the important protocols in use today—especially the Internet protocols. Peterson and Davie have a long involvement in and much experience with the Internet. Thus their book reflects not just the theoretical issues in x Foreword to the First Edition protocol design, but the real factors that matter in practice. The book looks at some of the protocols that are just emerging now, so the reader can be assured of an up-to-date perspective. But most importantly, the discussion of basic issues is presented in a way that derives from the fundamental nature of the problem, not the constraints of the layered reference model or the details of today’s protocols. In this regard, what this book presents is both timely and timeless. The combination of real-world relevance, current examples, and careful explanation of fundamentals makes this book unique. CONTENTS Foreword vii Foreword to the First Edition ix Preface xix 1 Foundation Problem: Building a Network 2 1.1 Applications 4 1.2 Requirements 6 1.2.1 Connectivity 7 1.2.2 Cost-Effective Resource Sharing 10 1.2.3 Support for Common Services 15 1.3 Network Architecture 19 1.3.1 Layering and Protocols 20 1.3.2 OSI Architecture 26 1.3.3 Internet Architecture 27 1.4 Implementing Network Software 30 1.4.1 Application Programming Interface (Sockets) 31 1.4.2 Example Application 33 1.4.3 Protocol Implementation Issues 36 1.5 Performance 40 1.5.1 Bandwidth and Latency 40 1.5.2 Delay × Bandwidth Product 44 1.5.3 High-Speed Networks 46 1.5.4 Application Performance Needs 48 1.6 Summary 50 Open Issue: Ubiquitous Networking 51 Further Reading 52 Exercises 55 xii Contents 2 Direct Link Networks Problem: Physically Connecting Hosts 64 2.1 Hardware Building Blocks 66 2.1.1 Nodes 66 2.1.2 Links 67 2.2 Encoding (NRZ, NRZI, Manchester, 4B/5B) 75 2.3 Framing 79 2.3.1 Byte-Oriented Protocols (BISYNC, PPP, DDCMP) 80 2.3.2 Bit-Oriented Protocols (HDLC) 83 2.3.3 Clock-Based Framing (SONET) 84 2.4 Error Detection 88 2.4.1 Two-Dimensional Parity 89 2.4.2 Internet Checksum Algorithm 90 2.4.3 Cyclic Redundancy Check 92 2.5 Reliable Transmission 97 2.5.1 Stop-and-Wait 98 2.5.2 Sliding Window 101 2.5.3 Concurrent Logical Channels 110 2.6 Ethernet (802.3) 111 2.6.1 Physical Properties 112 2.6.2 Access Protocol 114 2.6.3 Experience with Ethernet 119 2.7 Token Rings (802.5, FDDI) 120 2.7.1 Physical Properties 121 2.7.2 Token Ring Media Access Control 122 2.7.3 Token Ring Maintenance 125 2.7.4 Frame Format 126 2.7.5 FDDI 127 2.8 Wireless (802.11) 131 2.8.1 Physical Properties 132 2.8.2 Collision Avoidance 133 2.8.3 Distribution System 134 2.8.4 Frame Format 136 2.9 Network Adaptors 137 2.9.1 Components 138 2.9.2 View from the Host 139 2.9.3 Memory Bottleneck 144 2.10 Summary 146 Contents xiii Open Issue: Does It Belong in Hardware? 147 Further Reading 148 Exercises 150 3 Packet Switching Problem: Not All Networks Are Directly Connected 164 3.1 Switching and Forwarding 166 3.1.1 Datagrams 168 3.1.2 Virtual Circuit Switching 170 3.1.3 Source Routing 177 3.2 Bridges and LAN Switches 180 3.2.1 Learning Bridges 181 3.2.2 Spanning Tree Algorithm 185 3.2.3 Broadcast and Multicast 189 3.2.4 Limitations of Bridges 190 3.3 Cell Switching (ATM) 192 3.3.1 Cells 193 3.3.2 Segmentation and Reassembly 198 3.3.3 Virtual Paths 203 3.3.4 Physical Layers for ATM 204 3.3.5 ATM in the LAN 205 3.4 Implementation and Performance 210 3.4.1 Ports 212 3.4.2 Fabrics 216 3.5 Summary 220 Open Issue: The Future of ATM 221 Further Reading 221 Exercises 223 4 Internetworking Problem: There Is More Than One Network 234 4.1 Simple Internetworking (IP) 236 4.1.1 What Is an Internetwork? 236 4.1.2 Service Model 238 4.1.3 Global Addresses 250 4.1.4 Datagram Forwarding in IP 252 4.1.5 Address Translation (ARP) 258 4.1.6 Host Configuration (DHCP) 263 xiv Contents 4.1.7 Error Reporting (ICMP) 266 4.1.8 Virtual Networks and Tunnels 267 4.2 Routing 271 4.2.1 Network as a Graph 272 4.2.2 Distance Vector (RIP) 274 4.2.3 Link State (OSPF) 282 4.2.4 Metrics 291 4.2.5 Routing for Mobile Hosts 295 4.3 Global Internet 299 4.3.1 Subnetting 301 4.3.2 Classless Routing (CIDR) 306 4.3.3 Interdomain Routing (BGP) 308 4.3.4 Routing Areas 316 4.3.5 IP Version 6 (IPv6) 318 4.4 Multicast 331 4.4.1 Link-State Multicast 332 4.4.2 Distance-Vector Multicast 332 4.4.3 Protocol Independent Multicast (PIM) 336 4.5 Multiprotocol Label Switching (MPLS) 340 4.5.1 Destination-Based Forwarding 340 4.5.2 Explicit Routing 346 4.5.3 Virtual Private Networks and Tunnels 348 4.6 Summary 352 Open Issue: Deployment of IPV6 353 Further Reading 354 Exercises 355 5 End-to-End Protocols Problem: Getting Processess to Communicate 374 5.1 Simple Demultiplexer (UDP) 376 5.2 Reliable Byte Stream (TCP) 378 5.2.1 End-to-End Issues 379 5.2.2 Segment Format 382 5.2.3 Connection Establishment and Termination 384 5.2.4 Sliding Window Revisited 389 5.2.5 Triggering Transmission 395 5.2.6 Adaptive Retransmission 397 5.2.7 Record Boundaries 401 Contents xv 5.2.8 TCP Extensions 402 5.2.9 Alternative Design Choices 403 5.3 Remote Procedure Call 405 5.3.1 Bulk Transfer (BLAST) 408 5.3.2 Request/Reply (CHAN) 414 5.3.3 Dispatcher (SELECT) 423 5.3.4 Putting It All Together (SunRPC, DCE) 424 5.4 Performance 431 5.5 Summary 434 Open Issue: Application-Specific Protocols 435 Further Reading 436 Exercises 437 6 Congestion Control and Resource Allocation Problem: Allocating Resources 450 6.1 Issues in Resource Allocation 452 6.1.1 Network Model 452 6.1.2 Taxonomy 456 6.1.3 Evaluation Criteria 458 6.2 Queuing Disciplines 461 6.2.1 FIFO 462 6.2.2 Fair Queuing 463 6.3 TCP Congestion Control 468 6.3.1 Additive Increase/Multiplicative Decrease 468 6.3.2 Slow Start 471 6.3.3 Fast Retransmit and Fast Recovery 476 6.4 Congestion-Avoidance Mechanisms 478 6.4.1 DECbit 478 6.4.2 Random Early Detection (RED) 480 6.4.3 Source-Based Congestion Avoidance 486 6.5 Quality of Service 492 6.5.1 Application Requirements 493 6.5.2 Integrated Services (RSVP) 499 6.5.3 Differentiated Services (EF, AF) 509 6.5.4 ATM Quality of Service 513 6.5.5 Equation-Based Congestion Control 517 6.6 Summary 518 xvi Contents Open Issue: Inside versus Outside the Network 519 Further Reading 520 Exercises 521 7 End-to-End Data Problem: What Do We Do with the Data? 534 7.1 Presentation Formatting 536 7.1.1 Taxonomy 537 7.1.2 Examples (XDR, ASN.1, NDR) 541 7.1.3 Markup Languages (XML) 545 7.2 Data Compression 548 7.2.1 Lossless Compression Algorithms 550 7.2.2 Image Compression (JPEG) 552 7.2.3 Video Compression (MPEG) 557 7.2.4 Transmitting MPEG over a Network 562 7.2.5 Audio Compression (MP3) 566 7.3 Summary 568 Open Issue: Computer Networks Meet Consumer Electronics 568 Further Reading 569 Exercises 570 8 Network Security Problem: Securing the Data 578 8.1 Cryptographic Algorithms 580 8.1.1 Requirements 582 8.1.2 Secret Key Encryption (DES) 583 8.1.3 Public Key Encryption (RSA) 588 8.1.4 Message Digest Algorithms (MD5) 591 8.1.5 Implementation and Performance 593 8.2 Security Mechanisms 594 8.2.1 Authentication Protocols 594 8.2.2 Message Integrity Protocols 598 8.2.3 Public Key Distribution (X.509) 601 8.3 Example Systems 604 8.3.1 Pretty Good Privacy (PGP) 605 8.3.2 Secure Shell (SSH) 607 8.3.3 Transport Layer Security (TLS, SSL, HTTPS) 610 8.3.4 IP Security (IPSEC) 614 Contents xvii 8.4 Firewalls 617 8.4.1 Filter-Based Firewalls 618 8.4.2 Proxy-Based Firewalls 620 8.4.3 Limitations 622 8.5 Summary 622 Open Issue: Denial-of-Service Attacks 623 Further Reading 624 Exercises 625 9 Applications Problem: Applications Need Their Own Protocols 632 9.1 Name Service (DNS) 634 9.1.1 Domain Hierarchy 635 9.1.2 Name Servers 636 9.1.3 Name Resolution 640 9.2 Traditional Applications 643 9.2.1 Electronic Mail (SMTP, MIME, IMAP) 644 9.2.2 World Wide Web (HTTP) 650 9.2.3 Network Management (SNMP) 657 9.3 Multimedia Applications 660 9.3.1 Real-time Transport Protocol (RTP) 660 9.3.2 Session Control and Call Control (SDP, SIP, H.323) 671 9.4 Overlay Networks 680 9.4.1 Routing Overlays 682 9.4.2 Peer-to-Peer Networks 690 9.4.3 Content Distribution Networks 698 9.5 Summary 704 Open Issue: New Network Artichitecture 704 Further Reading 705 Exercises 707 Glossary 715 Bibliography 743 Solutions to Selected Exercises 763 Index 777 About the Authors 810 This Page Intentionally Left Blank PREFACE When the first edition of this book was published in 1996, it was a novelty to be able to order merchandise on the Internet, and a company that advertised its domain name was considered cutting edge. Today, Internet commerce is a fact of life, and “.com” stocks have gone through an entire boom and bust cycle. A host of new technologies ranging from optical switches to wireless networks are now becoming mainstream. It seems the only predictable thing about the Internet is constant change. Despite these changes the question we asked in the first edition is just as valid today: What are the underlying concepts and technologies that make the Internet work? The answer is that much of the TCP/IP architecture continues to function just as was envisioned by its creators nearly 30 years ago. This isn’t to say that the Internet architecture is uninteresting, quite the contrary. Understanding the design principles that underlie an architecture that has not only survived but fostered the kind of growth and change that the Internet has seen over the past three decades is precisely the right place to start. Like the previous editions, the third edition makes the “why” of the Internet architecture its cornerstone. Audience Our intent is that the book should serve as the text for a comprehensive networking class, at either the graduate or upper-division undergraduate level. We also believe that the book’s focus on core concepts should be appealing to industry professionals who are retraining for network-related assignments, as well as current network practitioners who want to understand the “whys” behind the protocols they work with every day and to see the big picture of networking. It is our experience that both students and professionals learning about networks for the first time often have the impression that network protocols are some sort of edict handed down from on high, and that their job is to learn as many TLAs (three-letter acronyms) as possible. In fact, protocols are the building blocks of a complex system developed through the application of engineering design principles. Moreover, they are constantly being refined, extended, and replaced based on real-world experience. xx Preface With this in mind, our goal with this book is to do more than survey the protocols in use today. Instead, we explain the underlying principles of sound network design. We feel that this grasp of underlying principles is the best tool for handling the rate of change in the networking field. Changes in the Third Edition Even though our focus is on the underlying principles of networking, we illustrate these principles using examples from today’s working Internet. Therefore, we added a significant amount of new material to track many of the important recent advances in networking. We also deleted, reorganized, and changed the focus of existing material to reflect changes that have taken place over the past seven years. Perhaps the most significant change we have noticed since writing the first edition is that almost every reader now has some familiarity with networked applications such as the World Wide Web and email. For this reason, we have increased the focus on applications, starting in the first chapter. We use applications as the motivation for the study of networking, and to derive a set of requirements that a useful network must meet if it is to support both current and future applications on a global scale. However, we retain the problem-solving approach of previous editions that starts with the problem of interconnecting hosts and works its way up the layers to conclude with a detailed examination of application-layer issues. We believe it is important to make the topics covered in the book relevant by starting with applications and their needs. At the same time, we feel that higher-layer issues, such as application-layer and transport- layer protocols, are best understood after the basic problems of connecting hosts and switching packets have been explained. Another important change in this edition is in the exercises. We have increased the number and quality of exercises; we have attempted to identify those that are especially difficult or that require above-average levels of mathematical knowledge (these are marked with an icon ); and in each chapter we have added a number of exercises with worked solutions that are included in the book. As before, the complete set of exercise solutions is available only to instructors. As we did in the second edition, we have added or increased coverage of im- portant new topics and brought other topics up-to-date. Major new or substantially updated topics in this edition are ■ a new section on Multiprotocol Label Switching (MPLS), including coverage of traffic engineering and virtual private networks ■ a new section on overlay networks, including “peer-to-peer” networking and “content distribution networks” Preface xxi ■ greatly expanded coverage on protocols for multimedia applications, such as Session Initiation Protocol (SIP) and Session Description Protocol (SDP) ■ updated coverage of congestion-control mechanisms, including selective ac- knowledgments for TCP, equation-based congestion control, and explicit con- gestion notification ■ updated security coverage, including distributed denial of service (DDoS) at- tacks ■ updated material on wireless technology, including spread spectrum tech- niques and the emerging 802.11 standards Finally, the book is now supplemented by a comprehensive set of laboratory exer- cises designed to illustrate the key concepts through simulation experiments. Sections that discuss material covered by the laboratory exercises are marked with the icon shown in the margin. Details on this new feature of the book appear below. Approach For an area that’s as dynamic and changing as computer networks, the most important thing a textbook can offer is perspective—to distinguish between what’s important and what’s not, and between what’s lasting and what’s superficial. Based on our experi- ence over the past 20 years doing research that has led to new networking technology, teaching undergraduate and graduate students about the latest trends in network- ing, and delivering advanced networking products to market, we have developed a perspective—which we call the systems approach—that forms the soul of this book. The systems approach has several implications: ■ Rather than accept existing artifacts as gospel, we start with first principles and walk you through the thought process that led to today’s networks. This allows us to explain why networks look like they do. It is our experience that once you understand the underlying concepts, any new protocol that you are confronted with will be relatively easy to digest. ■ Although the material is loosely organized around the traditional network layers, starting at the bottom and moving up the protocol stack, we do not adopt a rigid layered approach. Many topics—congestion control and security are good examples—have implications up and down the hierarchy, and so we discuss them outside the traditional layered model. In short, we believe layering makes a good servant but a poor master; it’s more often useful to take an end-to-end perspective. xxii Preface ■ Rather than explain how protocols work in the abstract, we use the most important protocols in use today—many of them from the TCP/IP Internet— to illustrate how networks work in practice. This allows us to include real- world experiences in the discussion. ■ Although at the lowest levels networks are constructed from commodity hard- ware that can be bought from computer vendors and communication services that can be leased from the phone company, it is the software that allows net- works to provide new services and adapt quickly to changing circumstances. It is for this reason that we emphasize how network software is implemented, rather than stopping with a description of the abstract algorithms involved. We also include code segments taken from a working protocol stack to illus- trate how you might implement certain protocols and algorithms. ■ Networks are constructed from many building-block pieces, and while it is necessary to be able to abstract away uninteresting elements when solving a particular problem, it is essential to understand how all the pieces fit to- gether to form a functioning network. We therefore spend considerable time explaining the overall end-to-end behavior of networks, not just the individ- ual components, so that it is possible to understand how a complete network operates, all the way from the application to the hardware. ■ The systems approach implies doing experimental performance studies, and then using the data you gather both to quantitatively analyze various design options and to guide you in optimizing the implementation. This emphasis on empirical analysis pervades the book. ■ Networks are like other computer systems—for example, operating systems, processor architectures, distributed and parallel systems, and so on. They are all large and complex. To help manage this complexity, system builders often draw on a collection of design principles. We highlight these design principles as they are introduced throughout the book, illustrated, of course, with examples from computer networks. Pedagogy and Features The third edition retains several features that we encourage you to take advantage of: ■ Problem statements. At the start of each chapter, we describe a problem that identifies the next set of issues that must be addressed in the design of a network. This statement introduces and motivates the issues to be explored in the chapter. Preface xxiii ■ Shaded sidebars. Throughout the text, shaded sidebars elaborate on the topic being discussed or introduce a related advanced topic. In many cases, these sidebars relate real-world anecdotes about networking. ■ Highlighted paragraphs. These paragraphs summarize an important nugget of information that we want you to take away from the discussion, such as a widely applicable system design principle. ■ Real protocols. Even though the book’s focus is on core concepts rather than existing protocol specifications, real protocols are used to illustrate most of the important ideas. As a result, the book can be used as a source of reference for many protocols. To help you find the descriptions of the protocols, each appli- cable section heading parenthetically identifies the protocols described in that section. For example, Section 5.2, which describes the principles of reliable end-to-end protocols, provides a detailed description of TCP, the canonical example of such a protocol. ■ Open issues. We conclude the main body of each chapter with an important issue that is currently being debated in the research community, the commer- cial world, or society as a whole. We have found that discussing these issues helps to make the subject of networking more relevant and exciting. ■ Further reading. These highly selective lists appear at the end of each chapter. Each list generally contains the seminal papers on the topics just discussed. We strongly recommend that advanced readers (e.g., graduate students) study the papers in this reading list to supplement the material covered in the chapter. Road Map and Course Use The book is organized as follows: ■ Chapter 1 introduces the set of core ideas that are used throughout the rest of the text. Motivated by widespread applications, it discusses what goes into network architecture, and it defines the quantitative performance metrics that often drive network design. ■ Chapter 2 surveys a wide range of low-level network technologies, ranging from Ethernet to token ring to wireless. It also describes many of the issues that all data link protocols must address, including encoding, framing, and error detection. xxiv Preface ■ Chapter 3 introduces the basic models of switched networks (datagrams versus virtual circuits) and describes one prevalent switching technology (ATM) in some detail. It also discusses the design of hardware-based switches. ■ Chapter 4 introduces internetworking and describes the key elements of the Internet Protocol (IP). A central question addressed in this chapter is how networks that scale to the size of the Internet are able to route packets. ■ Chapter 5 moves up to the transport level, describing both the Internet’s Trans- mission Control Protocol (TCP) and Remote Procedure Call (RPC) used to build client/server applications in detail. ■ Chapter 6 discusses congestion control and resource allocation. The issues in this chapter cut across both the network level (Chapters 3 and 4) and the transport level (Chapter 5). Of particular note, this chapter describes how congestion control works in TCP, and it introduces the mechanisms used by both the Internet and ATM to provide quality of service. ■ Chapter 7 considers the data sent through a network. This includes the prob- lems of both presentation formatting and data compression. The discussion of compression includes explanations of how MPEG video compression and MP3 audio compression work. ■ Chapter 8 discusses network security, ranging from an overview of cryptog- raphy protocols (DES, RSA, MD5), to protocols for security services (authen- tication, digital signature, message integrity), to complete security systems (privacy enhanced email, IPSEC). The chapter also discusses pragmatic issues like firewalls. ■ Chapter 9 describes a representative sample of network applications and the protocols they use, including traditional applications like email and the Web, multimedia applications such as IP telephony and video streaming, and overlay networks like peer-to-peer file sharing and content distribution networks. For an undergraduate course, extra class time will most likely be needed to help students digest the introductory material in the first chapter, probably at the expense of the more advanced topics covered in Chapters 6 through 8. Chapter 9 then returns to the popular topic of network applications. In contrast, the instructor for a graduate course should be able to cover the first chapter in only a lecture or two—with students studying the material more carefully on their own—thereby freeing up additional class time to cover the last four chapters in depth. Both graduate and undergraduate classes will want to cover the core material contained in the middle four chapters Preface xxv (Chapters 2–5), although an undergraduate class might choose to skim the more ad- vanced sections (e.g., Sections 2.2, 2.9, 3.4, and 4.4). For those of you using the book in self-study, we believe that the topics we have selected cover the core of computer networking, and so we recommend that the book be read sequentially, from front to back. In addition, we have included a liberal supply of references to help you locate supplementary material that is relevant to your specific areas of interest, and we have included solutions to selected exercises. The book takes a unique approach to the topic of congestion control by pulling all topics related to congestion control and resource allocation together in a single place—Chapter 6. We do this because the problem of congestion control cannot be solved at any one level, and we want you to consider the various design options at the same time. (This is consistent with our view that strict layering often obscures important design trade-offs.) A more traditional treatment of congestion control is possible, however, by studying Section 6.2 in the context of Chapter 3 and Section 6.3 in the context of Chapter 5. Exercises Significant effort has gone into improving the exercises in both the second and third editions. In the second edition we greatly increased the number of problems and, based on class testing, dramatically improved their quality. In this edition, we added a few more exercises, but made two other important changes: ■ For those exercises that we feel are particularly challenging or require special knowledge not provided in the book (e.g., probability expertise), we have added an icon to indicate the extra level of difficulty. ■ In each chapter we added some extra representative exercises for which worked solutions are provided in the back of the book. These exercises, marked , are intended to provide some help in tackling the other exercises in the book. The current sets of exercises are of several different styles: ■ Analytical exercises that ask the student to do simple algebraic calculations that demonstrate their understanding of fundamental relationships ■ Design questions that ask the student to propose and evaluate protocols for various circumstances ■ Hands-on questions that ask the student to write a few lines of code to test an idea or to experiment with an existing network utility xxvi Preface ■ Library research questions that ask the student to learn more about a partic- ular topic Also, as described in more detail below, socket-based programming assignments, as well as simulation labs, are available online. Supplemental Materials and Online Resources To assist instructors, we have prepared an instructor’s manual that contains solutions to selected exercises. The manual is available from the publisher. Additional support materials, including lecture slides, figures from the text, socket-based programming assignments, and sample exams and programming assign- ments are available through the Morgan Kaufmann Web site at http://www.mkp.com (search for Computer Networks). We suggest that you visit the page for this book every few weeks, as we will be adding support materials and establishing links to networking-related sites on a regular basis. And finally, new with the third edition, a set of laboratory experiments supple- ments the book. These labs, developed by Professor Emad Aboelela from the University of Massachusetts Dartmouth, use simulation to explore the behavior, scalability, and performance of protocols covered in the book. The simulations use the OPNET simu- lation toolset, which is available for free to anyone using Computer Networks in their course. Acknowledgments This book would not have been possible without the help of many people. We would like to thank them for their efforts in improving the end result. Before we do so, however, we should mention that we have done our best to correct the mistakes that the reviewers have pointed out and to accurately describe the protocols and mech- anisms that our colleagues have explained to us. We alone are responsible for any remaining errors. If you should find any of these, please send email to our publisher, Morgan Kaufmann, at netbugs@mkp.com, and we will endeavor to correct them in future printings of this book. First, we would like to thank the many people who reviewed drafts of all or parts of the manuscript. In addition to those who reviewed prior editions, we wish to thank Carl Emberger, Isaac Ghansah, and Bobby Bhattacharjee for their thor- ough reviews. Thanks also to Peter Druschel, Limin Wang, Aki Nakao, Dave Oran, George Swallow, Peter Lei, and Michael Ramalho for their reviews of various sec- tions. We also wish to thank all those who provided feedback and input to help us decide what to do in this edition: Chedley Aouriri, Peter Steenkiste, Esther A. Hughes, Ping-Tsai Chung, Doug Szajda, Mark Andersland, Leo Tam, C. P. Watkins, Preface xxvii Brian L. Mark, Miguel A. Labrador, Gene Chase, Harry W. Tyrer, Robert Siegfried, Harlan B. Russell, John R. Black, Robert Y. Ling, Julia Johnson, Karen Collins, Clark Verbrugge, Monjy Rabemanantsoa, Kerry D. LaViolette, William Honig, Kevin Mills, Murat Demirer, J Rufinus, Manton Matthews, Errin W. Fulp, Wayne Daniel, Luiz DaSilva, Don Yates, Raouf Boules, Nick McKeown, Neil T. Spring, Kris Verma, Szuecs Laszlo, Ted Herman, Mark Sternhagen, Zongming Fei, Dulal C. Kar, Mingyan Liu, Ken Surendran, Rakesh Arya, Mario J. Gonzalez, Annie Stanton, Tim Batten, and Paul Francis. Second, several members of the Network Systems Group at Princeton contributed ideas, examples, corrections, data, and code to this book. In particular, we would like to thank Andy Bavier, Tammo Spalink, Mike Wawrzoniak, Zuki Gottlieb, George Tzanetakis, and Chad Mynhier. As before, we want to thank the Defense Advanced Research Projects Agency, the National Science Foundation, Intel Corporation, and Cisco Systems, Inc. for supporting our networking research over the past several years. Third, we would like to thank our series editor, David Clark, as well as all the people at Morgan Kaufmann who helped shepherd us through the book-writing process. A special thanks is due to our original sponsoring editor, Jennifer Mann; our editor for the third edition, Rick Adams; our developmental editor, Karyn Johnson; and our production manager, Simon Crump. The whole crew at MKP has been a delight to work with. This Page Intentionally Left Blank Foundation I must Create a System, or be enslav’d by another Man’s; I will not Reason and Compare: my business is to Create. —William Blake Suppose you want to build a computer network, one that has the potential to grow to global proportions and to support applications as diverse as telecon- ferencing, video-on-demand, electronic commerce, distributed computing, and digital libraries. What available technologies would serve as the underlying build- ing blocks, and what kind of software architecture would you design to integrate PROBLEM Building a Network these building blocks into an effec- tive communication service? Answer- ing this question is the overriding goal of this book—to describe the available building materials and then to show how they can be used to con- struct a network from the ground up. Before we can understand how to design a computer network, we should first agree on exactly what a computer network is. At one time, the term network meant the set of serial lines used to attach dumb terminals to mainframe computers. To some, the term implies the voice telephone network. To others, the only interesting network is the cable network used to disseminate video signals. The main thing these networks have in common is that they are specialized to handle one particular kind of data (keystrokes, voice, or video) and they typically connect to special-purpose devices (terminals, hand receivers, and television sets). What distinguishes a computer network from these other types of networks? Probably the most important characteristic of a computer network is its generality. Computer networks are built primarily from general-purpose programmable hard- ware, and they are not optimized for a particular application like making phone calls or delivering television signals. Instead, they are able to carry many different types of data, and they support a wide, and ever-growing, range of applications. This chapter looks 1at some typical applications of computer networks and discusses the requirements that a network designer who wishes to support such applications must be aware of. Once we understand the requirements, how do we proceed? Fortunately, we will not be building the first net- work. Others, most notably the community of researchers responsible for the Internet, have gone before us. We will use the wealth of experience generated from the Internet to guide our design. This experience is embodied in a net- work architecture that identifies the available hardware and software components and shows how they can be arranged to form a complete network system. To start us on the road toward understanding how to build a network, this chapter does four things. First, it explores the requirements that different applications and different communities of people (such as network users and network operators) place on the network. Second, it introduces the idea of a network architecture, which lays the foundation for the rest of the book. Third, it intro- duces some of the key elements in the implementation of computer networks. Finally, it identifies the key metrics that are used to evaluate the performance of computer networks. 4 1 Foundation 1.1 Applications Most people know the Internet through its applications: the World Wide Web, email, streaming audio and video, chat rooms, and music (file) sharing. The Web, for example, presents an intuitively simple interface. Users view pages full of textual and graphical objects, click on objects that they want to learn more about, and a corresponding new page appears. Most people are also aware that just under the covers, each selectable object on a page is bound to an identifier for the next page to be viewed. This identifier, called a uniform resource locator (URL), uniquely names every possible page that can be viewed from your Web browser. For example, http://www.mkp.com/pd3e is the URL for a page representing this book at Morgan Kaufmann: The string http indicates that the HyperText Transfer Protocol (HTTP) should be used to download the page, www.mkp.com is the name of the machine that serves the page, and pd3e uniquely identifies the page at the publisher’s site. What most Web users are not aware of, however, is that by clicking on just one such URL, as many as 17 messages may be exchanged over the Internet, and this assumes the page itself is small enough to fit in a single message. This number includes up to six messages to translate the server name (www.mkp.com) into its Internet address (213.38.165.180), three messages to set up a Transmission Control Protocol (TCP) connection between your browser and this server, four messages for your browser to send the HTTP “get” request and the server to respond with the requested page (and for each side to acknowledge receipt of that message), and four messages to tear down the TCP connection. Of course, this does not include the millions of messages exchanged by Internet nodes throughout the day, just to let each other know that they exist and are ready to serve Web pages, translate names to addresses, and forward messages toward their ultimate destination. Although not yet as common as surfing the Web, another emerging application of the Internet is streaming audio and video. Although an entire video file could first be fetched from a remote machine and then played on the local machine, similar to the process of downloading and displaying a Web page, this would entail waiting for the last second of the video file to be delivered before starting to look at it. Streaming video implies that the sender and the receiver are, respectively, the source and the sink for the video stream. That is, the source generates a video stream (perhaps using a video capture card), sends it across the Internet in messages, and the sink displays the stream as it arrives. To be more precise, video is not an application; it is a type of data. One example of a video application is video-on-demand, which reads a preexisting movie from disk 1.1 Applications 5 and transmits it over the network. Another kind of application is videoconferencing, which is actually the more interesting case because it has very tight timing constraints. Just as when using the telephone, the interactions among the participants must be timely. When a person at one end gestures, then that action must be displayed at the other end as quickly as possible. Too much delay makes the system unusable. In contrast, if it takes several seconds from the time the user starts the video until the first image is displayed, then the service is still deemed satisfactory. Also, interactive video usually implies that video is flowing in both directions, while a video-on-demand application is most likely sending video in only one direction. The Unix application vic is an example of a popular videoconferencing tool. Figure 1.1 shows the control panel for a vic session. Note that vic is actually one of a suite of conferencing tools designed at Lawrence Berkeley Laboratory and Figure 1.1 The vic video application. 6 1 Foundation UC Berkeley. The others include a whiteboard application (wb) that allows users to send sketches and slides to each other, a visual audio tool called vat, and a session directory (sdr) that is used to create and advertise videoconferences. All these tools run on Unix—hence their lowercase names—and are freely available on the Internet. Similar tools are available for other operating systems. Although they are just two examples, downloading pages from the Web and participating in a videoconference demonstrate the diversity of applications that can be built on top of the Internet and hint at the complexity of the Internet’s design. Starting from the beginning, and addressing one problem at a time, the rest of this book explains how to build a network that supports such a wide range of applications. Chapter 9 concludes the book by revisiting these two specific applications, as well as several others that have become popular on today’s Internet. 1.2 Requirements We have just established an ambitious goal for ourselves: to understand how to build a computer network from the ground up. Our approach to accomplishing this goal will be to start from first principles, and then ask the kinds of questions we would naturally ask if building an actual network. At each step, we will use today’s proto- cols to illustrate various design choices available to us, but we will not accept these existing artifacts as gospel. Instead, we will be asking (and answering) the question of why networks are designed the way they are. While it is tempting to settle for just understanding the way it’s done today, it is important to recognize the underlying con- cepts because networks are constantly changing as the technology evolves and new applications are invented. It is our experience that once you understand the funda- mental ideas, any new protocol that you are confronted with will be relatively easy to digest. The first step is to identify the set of constraints and requirements that influence network design. Before getting started, however, it is important to understand that the expectations you have of a network depend on your perspective: ■ An application programmer would list the services that his or her application needs, for example, a guarantee that each message the application sends will be delivered without error within a certain amount of time. ■ A network designer would list the properties of a cost-effective design, for example, that network resources are efficiently utilized and fairly allocated to different users. ■ A network provider would list the characteristics of a system that is easy to administer and manage, for example, in which faults can be easily isolated and where it is easy to account for usage. 1.2 Requirements 7 This section attempts to distill these different perspectives into a high-level introduction to the major considerations that drive network design, and in doing so, identifies the challenges addressed throughout the rest of this book. 1.2.1 Connectivity Starting with the obvious, a network must provide connectivity among a set of com- puters. Sometimes it is enough to build a limited network that connects only a few select machines. In fact, for reasons of privacy and security, many private (corporate) networks have the explicit goal of limiting the set of machines that are connected. In contrast, other networks (of which the Internet is the prime example) are designed to grow in a way that allows them the potential to connect all the computers in the world. A system that is designed to support growth to an arbitrarily large size is said to scale. Using the Internet as a model, this book addresses the challenge of scalability. Links, Nodes, and Clouds Network connectivity occurs at many different levels. At the lowest level, a network can consist of two or more computers directly connected by some physical medium, such as a coaxial cable or an optical fiber. We call such a physical medium a link, and we often refer to the computers it connects as nodes. (Sometimes a node is a more specialized piece of hardware rather than a computer, but we overlook that distinction for the purposes of this discussion.) As illustrated in Figure 1.2, physical links are sometimes limited to a pair of nodes (such a link is said to be point-to-point), while in other cases, more than two nodes may share a single physical link (such a link is said to be multiple access). Whether a given link supports point-to-point or multiple- access connectivity depends on how the node is attached to the link. It is also the case that multiple-access links are often limited in size, in terms of both the geographical distance they can cover and the number of nodes they can connect. The exception is a satellite link, which can cover a wide geographic area. (a) (b) … Figure 1.2 Direct links: (a) point-to-point; (b) multiple-access. 8 1 Foundation Figure 1.3 Switched network. If computer networks were limited to situations in which all nodes are directly connected to each other over a common physical medium, then either networks would be very limited in the number of computers they could connect, or the number of wires coming out of the back of each node would quickly become both unmanageable and very expensive. Fortunately, connectivity between two nodes does not necessarily imply a direct physical connection between them—indirect connectivity may be achieved among a set of cooperating nodes. Consider the following two examples of how a collection of computers can be indirectly connected. Figure 1.3 shows a set of nodes, each of which is attached to one or more point- to-point links. Those nodes that are attached to at least two links run software that for- wards data received on one link out on another. If organized in a systematic way, these forwarding nodes form a switched network. There are numerous types of switched net- works, of which the two most common are circuit switched and packet switched. The former is most notably employed by the telephone system, while the latter is used for the overwhelming majority of computer networks and will be the focus of this book. The important feature of packet-switched networks is that the nodes in such a network send discrete blocks of data to each other. Think of these blocks of data as correspond- ing to some piece of application data such as a file, a piece of email, or an image. We call each block of data either a packet or a message, and for now we use these terms interchangeably; we discuss the reason they are not always the same in Section 1.2.2. Packet-switched networks typically use a strategy called store-and-forward.As the name suggests, each node in a store-and-forward network first receives a complete 1.2 Requirements 9 packet over some link, stores the packet in its internal memory, and then forwards the complete packet to the next node. In contrast, a circuit-switched network first establishes a dedicated circuit across a sequence of links and then allows the source node to send a stream of bits across this circuit to a destination node. The major reason for using packet switching rather than circuit switching in a computer network is efficiency, discussed in the next subsection. The cloud in Figure 1.3 distinguishes between the nodes on the inside that implement the network (they are commonly called switches, and their sole func- tion is to store and forward packets) and the nodes on the outside of the cloud that use the network (they are commonly called hosts, and they support users and run application programs). Also note that the cloud in Figure 1.3 is one of the most important icons of computer networking. In general, we use a cloud to denote any type of network, whether it is a single point-to-point link, a multiple-access link, or a switched network. Thus, whenever you see a cloud used in a figure, you can think of it as a placeholder for any of the networking technologies covered in this book. A second way in which a set of computers can be indirectly connected is shown in Figure 1.4. In this situation, a set of independent networks (clouds) are interconnected to form an internetwork, or internet for short. We adopt the Internet’s convention of referring to a generic internetwork of networks as a lowercase i internet, and the Figure 1.4 Interconnection of networks. 10 1 Foundation currently operational TCP/IP Internet as the capital I Internet. A node that is connected to two or more networks is commonly called a router or gateway, and it plays much the same role as a switch—it forwards messages from one network to another. Note that an internet can itself be viewed as another kind of network, which means that an internet can be built from an interconnection of internets. Thus, we can recursively build arbitrarily large networks by interconnecting clouds to form larger clouds. Just because a set of hosts are directly or indirectly connected to each other does not mean that we have succeeded in providing host-to-host connectivity. The final requirement is that each node must be able to say which of the other nodes on the network it wants to communicate with. This is done by assigning an address to each node. An address is a byte string that identifies a node; that is, the network can use a node’s address to distinguish it from the other nodes connected to the network. When a source node wants the network to deliver a message to a certain destination node, it specifies the address of the destination node. If the sending and receiving nodes are not directly connected, then the switches and routers of the network use this address to decide how to forward the message toward the destination. The process of determining systematically how to forward messages toward the destination node based on its address is called routing. This brief introduction to addressing and routing has presumed that the source node wants to send a message to a single destination node (unicast). While this is the most common scenario, it is also possible that the source node might want to broadcast a message to all the nodes on the network. Or a source node might want to send a message to some subset of the other nodes, but not all of them, a situation called multicast. Thus, in addition to node-specific addresses, another requirement of a network is that it support multicast and broadcast addresses.◮ The main idea to take away from this discussion is that we can define a network recursively as consisting of two or more nodes connected by a physical link, or as two or more networks connected by a node. In other words, a network can be constructed from a nesting of networks, where at the bottom level, the network is implemented by some physical medium. One of the key challenges in providing network connectivity is to define an address for each node that is reachable on the network (including support for broadcast and multicast connectivity), and to be able to use this address to route messages toward the appropriate destination node(s). 1.2.2 Cost-Effective Resource Sharing As stated above, this book focuses on packet-switched networks. This section explains the key requirement of computer networks—efficiency—that leads us to packet switch- ing as the strategy of choice. 1.2 Requirements 11 Given a collection of nodes indirectly connected by a nesting of networks, it is possible for any pair of hosts to send messages to each other across a sequence of links and nodes. Of course, we want to do more than support just one pair of com- municating hosts—we want to provide all pairs of hosts with the ability to exchange messages. The question then is, How do all the hosts that want to communicate share the network, especially if they want to use it at the same time? And, as if that problem isn’t hard enough, how do several hosts share the same link when they all want to use it at the same time? To understand how hosts share a network, we need to introduce a fundamental concept, multiplexing, which means that a system resource is shared among multiple users. At an intuitive level, multiplexing can be explained by analogy to a timesharing computer system, where a single physical CPU is shared (multiplexed) among multiple jobs, each of which believes it has its own private processor. Similarly, data being sent by multiple users can be multiplexed over the physical links that make up a network. To see how this might work, consider the simple network illustrated in Figure 1.5, where the three hosts on the left side of the network (L1–L3) are sending data to the three hosts on the right (R1–R3) by sharing a switched network that contains only one physical link. (For simplicity, assume that host L1 is communicating with host R1, and so on.) In this situation, three flows of data—corresponding to the three pairs of hosts—are multiplexed onto a single physical link by switch 1 and then demultiplexed back into separate flows by switch 2. Note that we are being intentionally vague about exactly what a “flow of data” corresponds to. For the purposes of this discussion, assume that each host on the left has a large supply of data that it wants to send to its counterpart on the right. There are several different methods for multiplexing multiple flows onto one phy- sical link. One common method is synchronous time-division multiplexing (STDM). The idea of STDM is to divide time into equal-sized quanta and, in a round-robin L2 L3 R2 R3 L1 R1 Switch 1 Switch 2 Figure 1.5 Multiplexing multiple logical flows over a single physical link. 12 1 Foundation fashion, give each flow a chance to send its data over the physical link. In other words, during time quantum 1, data from the first flow is transmitted; during time quantum 2, data from the second flow is transmitted; and so on. This process continues until all the flows have had a turn, at which time the first flow gets to go again, and the process repeats. Another method is frequency- division multiplexing (FDM). The idea of FDM is to transmit each flow over the phys- ical link at a different frequency, much the same way that the signals for different TV stations are transmitted at a different fre- quency on a physical cable TV link. Although simple to understand, both STDM and FDM are limited in two ways. First, if one of the flows (host pairs) does not have any data to send, its share of the phys- ical link—that is, its time quantum or its frequency—remains idle, even if one of the other flows has data to transmit. For com- puter communication, the amount of time that a link is idle can be very large—for example, consider the amount of time you spend reading a Web page (leaving the link idle) compared to the time you spend fetch- ing the page. Second, both STDM and FDM are limited to situations in which the max- imum number of flows is fixed and known ahead of time. It is not practical to resize the quantum or to add additional quanta in the case of STDM or to add new frequencies in the case of FDM. The form of multiplexing that we make most use of in this book is called statistical multiplexing. Although the name is not all that helpful for understanding the concept, statistical multiplexing is really quite simple, with two key ideas. First, it is like STDM in that the physical link is shared over time—first data from one flow is transmitted over the physical link, SANs, LANs, MANs, and WANs One way to characterize networks is according to their size. Two well- known examples are LANs (local area networks) and WANs (wide area networks); the former typi- cally extend less than 1 km, while the latter can be worldwide. Other networks are classified as MANs (metropolitan area networks), which usually span tens of kilome- ters. The reason such classifications are interesting is that the size of a network often has implications for the underlying technology that can be used, with a key factor being the amount of time it takes for data to propagate from one end of the network to the other; we discuss this issue more in later chapters. An interesting historical note is that the term “wide area network” was not applied to the first WANs because there was no other sort of network to differentiate them from. When computers were in- credibly rare and expensive, there was no point in thinking about how to connect all the computers in the local area—there was only one computer in that area. Only as 1.2 Requirements 13 then data from another flow is transmitted, and so on. Unlike STDM, however, data is transmitted from each flow on demand rather than during a predetermined time slot. Thus, if only one flow has data to send, it gets to transmit that data without waiting for its quantum to come around and thus without having to watch the quanta assigned computers began to proliferate did LANs become necessary, and the term “WAN” was then introduced to describe the larger networks that interconnected geographically dis- tant computers. Another kind of network that we need to be aware of is SANs (system area networks). SANs are usually confined to a single room and connect the various compo- nents of a large computing sys- tem. For example, HiPPI (High Performance Parallel Interface) and Fiber Channel are two common SAN technologies used to connect massively parallel processors to scalable storage servers and data vaults. (Because they often connect computers to storage servers, SANs are sometimes defined as storage area networks.) Although this book does not describe such networks in detail, they are worth knowing about because they are often at the leading edge in terms of perfor- mance, and because it is increas- ingly common to connect such net- works into LANs and WANs. to the other flows go by unused. It is this avoidance of idle time that gives packet switching its efficiency. As defined so far, however, statistical multiplexing has no mechanism to ensure that all the flows eventually get their turn to transmit over the physical link. That is, once a flow begins sending data, we need some way to limit the transmission, so that the other flows can have a turn. To account for this need, statistical multiplexing defines an upper bound on the size of the block of data that each flow is permitted to transmit at a given time. This limited-size block of data is typically referred to as a packet, to distin- guish it from the arbitrarily large message that an application program might want to transmit. Because a packet-switched net- work limits the maximum size of packets, a host may not be able to send a com- plete message in one packet. The source may need to fragment the message into several packets, with the receiver reassembling the packets back into the original message. In other words, each flow sends a se- quence of packets over the physical link, with a decision made on a packet-by-packet basis as to which flow’s packet to send next. Notice that if only one flow has data to send, then it can send a sequence of packets back- to-back. However, should more than one of the flows have data to send, then their pack- ets are interleaved on the link. Figure 1.6 depicts a switch multiplexing packets from multiple sources onto a single shared link. 14 1 Foundation … Figure 1.6 A switch multiplexing packets from multiple sources onto one shared link. The decision as to which packet to send next on a shared link can be made in a number of different ways. For example, in a network consisting of switches inter- connected by links such as the one in Figure 1.5, the decision would be made by the switch that transmits packets onto the shared link. (As we will see later, not all packet-switched networks actually involve switches, and they may use other mech- anisms to determine whose packet goes onto the link next.) Each switch in a packet- switched network makes this decision independently, on a packet-by-packet basis. One of the issues that faces a network designer is how to make this decision in a fair manner. For example, a switch could be designed to service packets on a first- in-first-out (FIFO) basis. Another approach would be to service the different flows in a round-robin manner, just as in STDM. This might be done to ensure that cer- tain flows receive a particular share of the link’s bandwidth, or that they never have their packets delayed in the switch for more than a certain length of time. A net- work that allows flows to request such treatment is said to support quality of service (QoS). Also, notice in Figure 1.6 that since the switch has to multiplex three incoming packet streams onto one outgoing link, it is possible that the switch will receive packets faster than the shared link can accommodate. In this case, the switch is forced to buffer these packets in its memory. Should a switch receive packets faster than it can send them for an extended period of time, then the switch will eventually run out of buffer space, and some packets will have to be dropped. When a switch is operating in this state, it is said to be congested. 1.2 Requirements 15 ◮ The bottom line is that statistical multiplexing defines a cost-effective way for multiple users (e.g., host-to-host flows of data) to share network resources (links and nodes) in a fine-grained manner. It defines the packet as the granularity with which the links of the network are allocated to different flows, with each switch able to schedule the use of the physical links it is connected to on a per-packet basis. Fairly allocating link capacity to different flows and dealing with congestion when it occurs are the key challenges of statistical multiplexing. 1.2.3 Support for Common Services While the previous section outlined the challenges involved in providing cost-effective connectivity among a group of hosts, it is overly simplistic to view a computer network as simply delivering packets among a collection of computers. It is more accurate to think of a network as providing the means for a set of application processes that are distributed over those computers to communicate. In other words, the next require- ment of a computer network is that the application programs running on the hosts connected to the network must be able to communicate in a meaningful way. When two application programs need to communicate with each other, there are a lot of complicated things that need to happen beyond simply sending a mes- sage from one host to another. One option would be for application designers to build all that complicated functionality into each application program. However, since many applications need common services, it is much more logical to implement those common services once and then to let the application designer build the application using those services. The challenge for a network designer is to identify the right set of common services. The goal is to hide the complexity of the network from the ap- plication without overly constraining the application designer. Intuitively, we view the network as providing logical channels over which application-level processes can communicate with each other; each channel provides the set of services required by that application. In other words, just as we use a cloud to abstractly represent connectivity among a set of computers, we now think of a chan- nel as connecting one process to another. Figure 1.7 shows a pair of application-level processes communicating over a logical channel that is, in turn, implemented on top of a cloud that connects a set of hosts. We can think of the channel as being like a pipe connecting two applications, so that a sending application can put data in one end and expect that data to be delivered by the network to the application at the other end of the pipe. The challenge is to recognize what functionality the channels should provide to application programs. For example, does the application require a guarantee that 16 1 Foundation Host HostHost Channel Application Host Application Host Figure 1.7 Processes communicating over an abstract channel. messages sent over the channel are delivered, or is it acceptable if some messages fail to arrive? Is it necessary that messages arrive at the recipient process in the same order in which they are sent, or does the recipient not care about the order in which messages arrive? Does the network need to ensure that no third parties are able to eavesdrop on the channel, or is privacy not a concern? In general, a network provides a variety of different types of channels, with each application selecting the type that best meets its needs. The rest of this section illustrates the thinking involved in defining useful channels. Identifying Common Communication Patterns Designing abstract channels involves first understanding the communication needs of a representative collection of applications, then extracting their common commu- nication requirements, and finally incorporating the functionality that meets these requirements in the network. One of the earliest applications supported on any network is a file access pro- gram like FTP (File Transfer Protocol) or NFS (Network File System). Although many details vary—for example, whether whole files are transferred across the network or only single blocks of the file are read/written at a given time—the communication 1.2 Requirements 17 component of remote file access is characterized by a pair of processes, one that re- quests that a file be read or written and a second process that honors this request. The process that requests access to the file is called the client, and the process that supports access to the file is called the server. Reading a file involves the client sending a small request message to a server and the server responding with a large message that contains the data in the file. Writing works in the opposite way—the client sends a large message containing the data to be written to the server, and the server responds with a small message confirming that the write to disk has taken place. A digital library, as exemplified by the World Wide Web, is another application that behaves in a similar way: A client process makes a request, and a server process responds by returning the requested data. Using file access, a digital library, and the two video applications described in the introduction (videoconferencing and video-on-demand) as a representative sample, we might decide to provide the following two types of channels: request/reply channels and message stream channels. The request/reply channel would be used by the file transfer and digital library applications. It would guarantee that every message sent by one side is received by the other side and that only one copy of each message is delivered. The request/reply channel might also protect the privacy and integrity of the data that flows over it, so that unauthorized parties cannot read or modify the data being exchanged between the client and server processes. The message stream channel could be used by both the video-on-demand and videoconferencing applications, provided it is parameterized to support both one-way and two-way traffic and to support different delay properties. The message stream channel might not need to guarantee that all messages are delivered, since a video appli- cation can operate adequately even if some frames are not received. It would, however, need to ensure that those messages that are delivered arrive in the same order in which they were sent, to avoid displaying frames out of sequence. Like the request/reply channel, the message stream channel might want to ensure the privacy and integrity of the video data. Finally, the message stream channel might need to support multicast, so that multiple parties can participate in the teleconference or view the video. While it is common for a network designer to strive for the smallest number of abstract channel types that can serve the largest number of applications, there is a danger in trying to get away with too few channel abstractions. Simply stated, if you have a hammer, then everything looks like a nail. For example, if all you have are message stream and request/reply channels, then it is tempting to use them for the next application that comes along, even if neither type provides exactly the semantics needed by the application. Thus, network designers will probably be inventing new types of channels—and adding options to existing channels—for as long as application programmers are inventing new applications. 18 1 Foundation Also note that independent of exactly what functionality a given channel pro- vides, there is the question of where that functionality is implemented. In many cases, it is easiest to view the host-to-host connectivity of the underlying network as simply providing a bit pipe, with any high-level communication semantics provided at the end hosts. The advantage of this approach is it keeps the switches in the middle of the network as simple as possible—they simply forward packets—but it requires the end hosts to take on much of the burden of supporting semantically rich process-to-process channels. The alternative is to push additional functionality onto the switches, thereby allowing the end hosts to be “dumb” devices (e.g., telephone handsets). We will see this question of how various network services are partitioned between the packet switches and the end hosts (devices) as a reoccurring issue in network design. Reliability As suggested by the examples just considered, reliable message delivery is one of the most important functions that a network can provide. It is difficult to determine how to provide this reliability, however, without first understanding how networks can fail. The first thing to recognize is that computer networks do not exist in a perfect world. Machines crash and later are rebooted, fibers are cut, electrical interference corrupts bits in the data being transmitted, switches run out of buffer space, and if these sorts of physical problems aren’t enough to worry about, the software that manages the hardware sometimes forwards packets into oblivion. Thus, a major requirement of a network is to mask (hide) certain kinds of failures, so as to make the network appear more reliable than it really is to the application programs using it. There are three general classes of failure that network designers have to worry about. First, as a packet is transmitted over a physical link, bit errors may be introduced into the data; that is,a1isturned intoa0orvice versa. Sometimes single bits are corrupted, but more often than not, a burst error occurs—several consecutive bits are corrupted. Bit errors typically occur because outside forces, such as lightning strikes, power surges, and microwave ovens, interfere with the transmission of data. The good news is that such bit errors are fairly rare, affecting on average only one out of every 106 to 107 bits on a typical copper-based cable and one out of every 1012 to 1014 bits on a typical optical fiber. As we will see, there are techniques that detect these bit errors with high probability. Once detected, it is sometimes possible to correct for such errors—if we know which bit or bits are corrupted, we can simply flip them—while in other cases the damage is so bad that it is necessary to discard the entire packet. In such a case, the sender may be expected to retransmit the packet. The second class of failure is at the packet, rather than the bit, level; that is, a complete packet is lost by the network. One reason this can happen is that the packet contains an uncorrectable bit error and therefore has to be discarded. A more likely 1.3 Network Architecture 19 reason, however, is that one of the nodes that has to handle the packet—for example, a switch that is forwarding it from one link to another—is so overloaded that it has no place to store the packet, and therefore is forced to drop it. This is the problem of congestion mentioned in Section 1.2.2. Less commonly, the software running on one of the nodes that handles the packet makes a mistake. For example, it might incorrectly forward a packet out on the wrong link, so that the packet never finds its way to the ultimate destination. As we will see, one of the main difficulties in dealing with lost packets is distinguishing between a packet that is indeed lost and one that is merely late in arriving at the destination. The third class of failure is at the node and link level; that is, a physical link is cut, or the computer it is connected to crashes. This can be caused by software that crashes, a power failure, or a reckless backhoe operator. While such failures can eventually be corrected, they can have a dramatic effect on the network for an extended period of time. However, they need not totally disable the network. In a packet-switched net- work, for example, it is sometimes possible to route around a failed node or link. One of the difficulties in dealing with this third class of failure is distinguishing between a failed computer and one that is merely slow, or in the case of a link, between one that has been cut and one that is very flaky and therefore introducing a high number of bit errors.◮ The key idea to take away from this discussion is that defining useful chan- nels involves both understanding the applications’ requirements and recognizing the limitations of the underlying technology. The challenge is to fill in the gap between what the application expects and what the underlying technology can provide. This is sometimes called the semantic gap. 1.3 Network Architecture In case you hadn’t noticed, the previous section established a pretty substantial set of requirements for network design—a computer network must provide general, cost- effective, fair, and robust connectivity among a large number of computers. As if this weren’t enough, networks do not remain fixed at any single point in time, but must evolve to accommodate changes in both the underlying technologies upon which they are based as well as changes in the demands placed on them by application programs. Designing a network to meet these requirements is no small task. To help deal with this complexity, network designers have developed general blueprints—usually called a network architecture—that guide the design and imple- mentation of networks. This section defines more carefully what we mean by a network architecture by introducing the central ideas that are common to all network archi- tectures. It also introduces two of the most widely referenced architectures—the OSI architecture and the Internet architecture. 20 1 Foundation Application programs Process-to-process channels Host-to-host connectivity Hardware Figure 1.8 Example of a layered network system. 1.3.1 Layering and Protocols When the system gets complex, the system designer introduces another level of ab- straction. The idea of an abstraction is to define a unifying model that can capture some important aspect of the system, encapsulate this model in an object that provides an interface that can be manipulated by other components of the system, and hide the details of how the object is implemented from the users of the object. The challenge is to identify abstractions that simultaneously provide a service that proves useful in a large number of situations and that can be efficiently implemented in the underlying system. This is exactly what we were doing when we introduced the idea of a channel in the previous section: We were providing an abstraction for applications that hides the complexity of the network from application writers. Abstractions naturally lead to layering, especially in network systems. The gen- eral idea is that you start with the services offered by the underlying hardware, and then add a sequence of layers, each providing a higher (more abstract) level of ser- vice. The services provided at the high layers are implemented in terms of the services provided by the low layers. Drawing on the discussion of requirements given in the previous section, for example, we might imagine a network as having two layers of abstraction sandwiched between the application program and the underlying hard- ware, as illustrated in Figure 1.8. The layer immediately above the hardware in this case might provide host-to-host connectivity, abstracting away the fact that there may be an arbitrarily complex network topology between any two hosts. The next layer up builds on the available host-to-host communication service and provides support for process-to-process channels, abstracting away the fact that the network occasionally loses messages, for example. Layering provides two nice features. First, it decomposes the problem of building a network into more manageable components. Rather than implementing a monolithic piece of software that does everything you will ever want, you can implement several layers, each of which solves one part of the problem. Second, it provides a more mod- ular design. If you decide that you want to add some new service, you may only need 1.3 Network Architecture 21 Hardware Host-to-host connectivity Request/reply channel Message stream channel Application programs Figure 1.9 Layered system with alternative abstractions available at a given layer. to modify the functionality at one layer, reusing the functions provided at all the other layers. Thinking of a system as a linear sequence of layers is an oversimplification, however. Many times there are multiple abstractions provided at any given level of the system, each providing a different service to the higher layers but building on the same low-level abstractions. To see this, consider the two types of channels discussed in Section 1.2.3: One provides a request/reply service, and one supports a message stream service. These two channels might be alternative offerings at some level of a multilevel networking system, as illustrated in Figure 1.9. Using this discussion of layering as a foundation, we are now ready to discuss the architecture of a network more precisely. For starters, the abstract objects that make up the layers of a network system are called protocols. That is, a protocol provides a communication service that higher-level objects (such as application processes, or perhaps higher-level protocols) use to exchange messages. For example, we could imag- ine a network that supports a request/reply protocol and a message stream protocol, corresponding to the request/reply and message stream channels discussed above. Each protocol defines two different interfaces. First, it defines a service inter- face to the other objects on the same computer that want to use its communication services. This service interface defines the operations that local objects can perform on the protocol. For example, a request/reply protocol would support operations by which an application can send and receive messages. Second, a protocol defines a peer interface to its counterpart (peer) on another machine. This second interface defines the form and meaning of messages exchanged between protocol peers to implement the communication service. This would determine the way in which a request/reply protocol on one machine communicates with its peer on another machine. In other words, a protocol defines a communication service that it exports locally, along with a set of rules governing the messages that the protocol exchanges with its peer(s) to implement this service. This situation is illustrated in Figure 1.10. 22 1 Foundation Host 1 Host 2 Service interface Peer-to-peer interface High-level object High-level object Protocol Protocol Figure 1.10 Service and peer interfaces. Except at the hardware level where peers directly communicate with each other over a link, peer-to-peer communication is indirect—each protocol communicates with its peer by passing messages to some lower-level protocol, which in turn delivers the message to its peer. In addition, there are potentially multiple protocols at any given level, each providing a different communication service. We therefore represent the suite of protocols that make up a network system with a protocol graph. The nodes of the graph correspond to protocols, and the edges represent a depends-on relation. For example, Figure 1.11 illustrates a protocol graph for the hypothetical layered system we have been discussing—protocols RRP (Request/Reply Protocol) and MSP (Mes- sage Stream Protocol) implement two different types of process-to-process channels, and both depend on HHP (Host-to-Host Protocol), which provides a host-to-host connectivity service. In this example, suppose that the file access program on host 1 wants to send a message to its peer on host 2 using the communication service offered by protocol RRP. In this case, the file application asks RRP to send the message on its behalf. To communicate with its peer, RRP then invokes the services of HHP, which in turn transmits the message to its peer on the other machine. Once the message has arrived at protocol HHP on host 2, HHP passes the message up to RRP, which in turn delivers the message to the file application. In this particular case, the application is said to employ the services of the protocol stack RRP/HHP. Note that the term protocol is used in two different ways. Sometimes it refers to the abstract interfaces—that is, the operations defined by the service interface and the form and meaning of messages exchanged between peers—and sometimes it refers to the module that actually implements these two interfaces. To distinguish between the 1.3 Network Architecture 23 Host 1 RRP MSP HHP Host 2 RRP MSP HHP File application Digital library application Video application File application Digital library application Video application Figure 1.11 Example of a protocol graph. interfaces and the module that implements these interfaces, we generally refer to the former as a protocol specification. Specifications are generally expressed using a combi- nation of prose, pseudocode, state transition diagrams, pictures of packet formats, and other abstract notations. It should be the case that a given protocol can be implemented in different ways by different programmers, as long as each adheres to the specification. The challenge is ensuring that two different implementations of the same specification can successfully exchange messages. Two or more protocol modules that do accurately implement a protocol specification are said to interoperate with each other. We can imagine many different protocols and protocol graphs that satisfy the communication requirements of a collection of applications. Fortunately, there exist standardization bodies, such as the International Standards Organization (ISO) and the Internet Engineering Task Force (IETF), that establish policies for a particular pro- tocol graph. We call the set of rules governing the form and content of a protocol graph a network architecture. Although beyond the scope of this book, standardiza- tion bodies such as the ISO and the IETF have established well-defined procedures for 24 1 Foundation introducing, validating, and finally approving protocols in their respective architec- tures. We briefly describe the architectures defined by the ISO and the IETF shortly, but first there are two additional things we need to explain about the mechanics of a protocol graph. Encapsulation Consider what happens in Figure 1.11 when one of the application programs sends a message to its peer by passing the message to protocol RRP. From RRP’s perspective, the message it is given by the application is an uninterpreted string of bytes. RRP does not care that these bytes represent an array of integers, an email message, a digital image, or whatever; it is simply charged with sending them to its peer. However, RRP must communicate control information to its peer, instructing it how to handle the message when it is received. RRP does this by attaching a header to the message. Generally speaking, a header is a small data structure—from a few bytes to a few dozen bytes—that is used among peers to communicate with each other. As the name suggests, headers are usually attached to the front of a message. In some cases, however, this peer-to-peer control information is sent at the end of the message, in which case it is called a trailer. The exact format for the header attached by RRP is defined by its protocol specification. The rest of the message—that is, the data being transmitted on behalf of the application—is called the message’s body or payload. We say that the application’s data is encapsulated in the new message created by protocol RRP. This process of encapsulation is then repeated at each level of the protocol graph; for example, HHP encapsulates RRP’s message by attaching a header of its own. If we now assume that HHP sends the message to its peer over some network, then when the message arrives at the destination host, it is processed in the opposite order: HHP first strips its header off the front of the message, interprets it (i.e., takes whatever action is appropriate given the contents of the header), and passes the body of the message up to RRP, which removes the header that its peer attached, takes whatever action is indicated by that header, and passes the body of the message up to the application program. The message passed up from RRP to the application on host 2 is exactly the same message as the application passed down to RRP on host 1; the application does not see any of the headers that have been attached to it to implement the lower-level communication services. This whole process is illustrated in Figure 1.12. Note that in this example, nodes in the network (e.g., switches and routers) may inspect the HHP header at the front of the message. Note that when we say a low-level protocol does not interpret the message it is given by some high-level protocol, we mean that it does not know how to extract any meaning from the data contained in the message. It is sometimes the case, however, that the low-level protocol applies some simple transformation to the data it is given, such as to compress or encrypt it. In this case, the protocol is transforming the entire 1.3 Network Architecture 25 Host 1 Host 2 Application program Application program RRP Data Data HHP RRP HHP Application program Application program RRP Data RRP Data HHP RRP Data Figure 1.12 High-level messages are encapsulated inside of low-level messages. body of the message, including both the original application’s data and all the headers attached to that data by higher-level protocols. Multiplexing and Demultiplexing Recall from Section 1.2.2 that a fundamental idea of packet switching is to multiplex multiple flows of data over a single physical link. This same idea applies up and down the protocol graph, not just to switching nodes. In Figure 1.11, for example, we can think of RRP as implementing a logical communication channel, with messages from two different applications multiplexed over this channel at the source host and then demultiplexed back to the appropriate application at the destination host. Practically speaking, all this means is that the header that RRP attaches to its messages contains an identifier that records the application to which the message belongs. We call this identifier RRP’s demultiplexing key,ordemux key for short. At the source host, RRP includes the appropriate demux key in its header. When the message is delivered to RRP on the destination host, it strips its header, examines the demux key, and demultiplexes the message to the correct application. 26 1 Foundation RRP is not unique in its support for multiplexing; nearly every protocol imple- ments this mechanism. For example, HHP has its own demux key to determine which messages to pass up to RRP and which to pass up to MSP.However, there is no uniform agreement among protocols—even those within a single network architecture—on ex- actly what constitutes a demux key. Some protocols use an 8-bit field (meaning they can support only 256 high-level protocols), and others use 16- or 32-bit fields. Also, some protocols have a single demultiplexing field in their header, while others have a pair of demultiplexing fields. In the former case, the same demux key is used on both sides of the communication, while in the latter case, each side uses a different key to identify the high-level protocol (or application program) to which the message is to be delivered. 1.3.2 OSI Architecture The ISO was one of the first organizations to formally define a common way to connect computers. Their architecture, called the Open Systems Interconnection (OSI) architec- ture and illustrated in Figure 1.13, defines a partitioning of network functionality into One or more nodes within the network End host Application Presentation Session Transport Network Data link Physical Network Data link Physical Network Data link Physical End host Application Presentation Session Transport Network Data link Physical Figure 1.13 OSI network architecture. 1.3 Network Architecture 27 seven layers, where one or more protocols implement the functionality assigned to a given layer. In this sense, the schematic given in Figure 1.13 is not a protocol graph, per se, but rather a reference model for a protocol graph. The ISO, usually in conjunction with a second standards organization known as the International Telecommunications Union (ITU),1 publishes a series of protocol specifications based on the OSI archi- tecture. This series is sometimes called the “X dot” series since the protocols are given names like X.25, X.400, X.500, and so on. There have been several networks based on these standards, including the public X.25 network and private networks like Tymnet. Starting at the bottom and working up, the physical layer handles the transmis- sion of raw bits over a communications link. The data link layer then collects a stream of bits into a larger aggregate called a frame. Network adaptors, along with device drivers running in the node’s OS, typically implement the data link level. This means that frames, not raw bits, are actually delivered to hosts. The network layer handles routing among nodes within a packet-switched network. At this layer, the unit of data exchanged among nodes is typically called a packet rather than a frame, although they are fundamentally the same thing. The lower three layers are implemented on all network nodes, including switches within the network and hosts connected along the exterior of the network. The transport layer then implements what we have up to this point been calling a process-to-process channel. Here, the unit of data exchanged is commonly called a message rather than a packet or a frame. The transport layer and higher layers typically run only on the end hosts and not on the intermediate switches or routers. There is less agreement about the definition of the top three layers. Skipping ahead to the top (seventh) layer, we find the application layer. Application layer pro- tocols include things like the File Transfer Protocol (FTP), which defines a protocol by which file transfer applications can interoperate. Below that, the presentation layer is concerned with the format of data exchanged between peers, for example, whether an integer is 16, 32, or 64 bits long and whether the most significant bit is transmitted first or last, or how a video stream is formatted. Finally, the session layer provides a name space that is used to tie together the potentially different transport streams that are part of a single application. For example, it might manage an audio stream and a video stream that are being combined in a teleconferencing application. 1.3.3 Internet Architecture The Internet architecture, which is also sometimes called the TCP/IP architecture after its two main protocols, is depicted in Figure 1.14. An alternative representation is given in Figure 1.15. The Internet architecture evolved out of experiences with an earlier 1A subcommittee of the ITU on telecommunications (ITU-T) replaces an earlier subcommittee of the ITU, which was known by its French name, Comit´e Consultatif International de T´el´egraphique et T´el´ephonique (CCITT). 28 1 Foundation … FTP TCP UDP IP NET1 NET2 NETn HTTP NV TFTP Figure 1.14 Internet protocol graph. TCP UDP IP Network Application Figure 1.15 Alternative view of the Internet architecture. packet-switched network called the ARPANET. Both the Internet and the ARPANET were funded by the Advanced Research Projects Agency (ARPA), one of the R&D funding agencies of the U.S. Department of Defense. The Internet and ARPANET were around before the OSI architecture, and the experience gained from building them was a major influence on the OSI reference model. While the seven-layer OSI model can, with some imagination, be applied to the Internet, a four-layer model is often used instead. At the lowest level are a wide variety of network protocols, denoted NET1, NET2, and so on. In practice, these protocols are implemented by a combination of hardware (e.g., a network adaptor) and soft- ware (e.g., a network device driver). For example, you might find Ethernet or Fiber Distributed Data Interface (FDDI) protocols at this layer. (These protocols in turn may actually involve several sublayers, but the Internet architecture does not presume anything about them.) The second layer consists of a single protocol—the Internet Protocol (IP). This is the protocol that supports the interconnection of multiple net- working technologies into a single, logical internetwork. The third layer contains two main protocols—the Transmission Control Protocol (TCP) and the User Datagram Protocol (UDP). TCP and UDP provide alternative logical channels to application 1.3 Network Architecture 29 programs: TCP provides a reliable byte-stream channel, and UDP provides an un- reliable datagram delivery channel (datagram may be thought of as a synonym for message). In the language of the Internet, TCP and UDP are sometimes called end-to- end protocols, although it is equally correct to refer to them as transport protocols. Running above the transport layer are a range of application protocols, such as FTP, TFTP (Trivial File Transport Protocol), Telnet (remote login), and SMTP (Simple Mail Transfer Protocol, or electronic mail), that enable the interoperation of popular applications. To understand the difference between an application layer protocol and an application, think of all the different World Wide Web browsers that are available (e.g., Mosaic, Netscape, Internet Explorer, Lynx, etc.). There are a similarly large number of different implementations of Web servers. The reason that you can use any one of these application programs to access a particular site on the Web is because they all conform to the same application layer protocol: HTTP (HyperText Transport Protocol). Confusingly, the same word sometimes applies to both an application and the application layer protocol that it uses (e.g., FTP). The Internet architecture has three features that are worth highlighting. First, as best illustrated by Figure 1.15, the Internet architecture does not imply strict layering. The application is free to bypass the defined transport layers and to directly use IP or one of the underlying networks. In fact, programmers are free to define new channel abstractions or applications that run on top of any of the existing protocols. Second, if you look closely at the protocol graph in Figure 1.14, you will notice an hourglass shape—wide at the top, narrow in the middle, and wide at the bottom. This shape actually reflects the central philosophy of the architecture. That is, IP serves as the focal point for the architecture—it defines a common method for exchanging packets among a wide collection of networks. Above IP can be arbitrarily many trans- port protocols, each offering a different channel abstraction to application programs. Thus, the issue of delivering messages from host to host is completely separated from the issue of providing a useful process-to-process communication service. Below IP, the architecture allows for arbitrarily many different network technologies, ranging from Ethernet to FDDI to ATM to single point-to-point links. A final attribute of the Internet architecture (or more accurately, of the IETF culture) is that in order for someone to propose a new protocol to be included in the architecture, they must produce both a protocol specification and at least one (and preferably two) representative implementations of the specification. The existence of working implementations is required for standards to be adopted by the IETF. This cultural assumption of the design community helps to ensure that the architecture’s protocols can be efficiently implemented. Perhaps the value the Internet culture places on working software is best exemplified by a quote on T-shirts commonly worn at IETF meetings: 30 1 Foundation We reject kings, presidents, and voting. We believe in rough consensus and running code. (Dave Clark) ◮ Of these three attributes of the Internet architecture, the hourglass design philos- ophy is important enough to bear repeating. The hourglass’s narrow waist represents a minimal and carefully chosen set of global capabilities that allows both higher-level applications and lower-level communication technologies to coexist, share capabili- ties, and evolve rapidly. The narrow-waisted model is critical to the Internet’s ability to adapt rapidly to new user demands and changing technologies. 1.4 Implementing Network Software Network architectures and protocol specifications are essential things, but a good blueprint is not enough to explain the phenomenal success of the Internet: The number of computers connected to the Internet has been doubling every year since 1981 and is now approaching 200 million; the number of people who use the Internet is estimated at well over 600 million; and it is believed that the number of bits transmitted over the Internet surpassed the corresponding figure for the voice phone system sometime in 2001. What explains the success of the Internet? There are certainly many contributing factors (including a good architecture), but one thing that has made the Internet such a runaway success is the fact that so much of its functionality is provided by software running in general-purpose computers. The significance of this is that new function- ality can be added readily with “just a small matter of programming.” As a result, new applications and services—electronic commerce, videoconferencing, and packet telephony, to name a few—have been showing up at a phenomenal pace. A related factor is the massive increase in computing power available in commod- ity machines. Although computer networks have always been capable in principle of transporting any kind of information, such as digital voice samples, digitized images, and so on, this potential was not particularly interesting if the computers sending and receiving that data were too slow to do anything useful with the information. Virtually all of today’s computers are capable of playing back digitized voice at full speed and can display video at a speed and resolution that is useful for some (but by no means all) applications. Thus, today’s networks have begun to support multimedia, and their support for it will only improve as computing hardware becomes faster. The point to take away from this is that knowing how to implement network software is an essential part of understanding computer networks. With this in mind, this section first introduces some of the issues involved in implementing an application program on top of a network, and then goes on to identify the issues involved in 1.4 Implementing Network Software 31 implementing the protocols running within the network. In many respects, network applications and network protocols are very similar—the way an application engages the services of the network is pretty much the same as the way a high-level protocol invokes the services of a low-level protocol. As we will see later in the section, however, there are a couple of important differences. 1.4.1 Application Programming Interface (Sockets) The place to start when implementing a network application is the interface exported by the network. Since most network protocols are implemented in software (espe- cially those high in the protocol stack), and nearly all computer systems implement their network protocols as part of the operating system, when we refer to the inter- face “exported by the network,” we are generally referring to the interface that the OS provides to its networking subsystem. This interface is often called the network application programming interface (API). Although each operating system is free to define its own network API (and most have), over time certain of these APIs have become widely supported; that is, they have been ported to operating systems other than their native system. This is what has happened with the socket interface originally provided by the Berkeley distribution of Unix, which is now supported in virtually all popular operating systems. The advantage of industrywide support for a single API is that applications can be easily ported from one OS to another. It is important to keep in mind, however, that application programs typically interact with many parts of the OS other than the network; for example, they read and write files, fork concurrent processes, and output to the graphical display. Just because two systems support the same network API does not mean that their file system, process, or graphic interfaces are the same. Still, understanding a widely adopted API like Unix sockets gives us a good place to start. Before describing the socket interface, it is important to keep two concerns sep- arate in your mind. Each protocol provides a certain set of services, and the API provides a syntax by which those services can be invoked in this particular OS. The implementation is then responsible for mapping the tangible set of operations and objects defined by the API onto the abstract set of services defined by the pro- tocol. If you have done a good job of defining the interface, then it will be possible to use the syntax of the interface to invoke the services of many different protocols. Such generality was certainly a goal of the socket interface, although it’s far from perfect. The main abstraction of the socket interface, not surprisingly, is the socket.A good way to think of a socket is as the point where a local application process attaches to the network. The interface defines operations for creating a socket, attaching the socket to the network, sending/receiving messages through the socket, and closing the 32 1 Foundation socket. To simplify the discussion, we will limit ourselves to showing how sockets are used with TCP. The first step is to create a socket, which is done with the following operation: int socket(int domain, int type, int protocol) The reason that this operation takes three arguments is that the socket interface was designed to be general enough to support any underlying protocol suite. Specifically, the domain argument specifies the protocol family that is going to be used: PF INET denotes the Internet family, PF UNIX denotes the Unix pipe facility, and PF PACKET denotes direct access to the network interface (i.e., it bypasses the TCP/IP protocol stack). The type argument indicates the semantics of the communication. SOCK STREAM is used to denote a byte stream. SOCK DGRAM is an alternative that denotes a message-oriented service, such as that provided by UDP. The protocol argument identifies the specific protocol that is going to be used. In our case, this argument is UNSPEC because the combination of PF INET and SOCK STREAM implies TCP.Finally, the return value from socket is a handle for the newly created socket, that is, an identifier by which we can refer to the socket in the future. It is given as an argument to subsequent operations on this socket. The next step depends on whether you are a client or a server. On a server machine, the application process performs a passive open—the server says that it is prepared to accept connections, but it does not actually establish a connection. The server does this by invoking the following three operations: int bind(int socket, struct sockaddr *address, int addr len) int listen(int socket, int backlog) int accept(int socket, struct sockaddr *address, int *addr len) The bind operation, as its name suggests, binds the newly created socket to the specified address. This is the network address of the local participant—the server. Note that, when used with the Internet protocols, address is a data structure that includes both the IP address of the server and a TCP port number. (As we will see in Chapter 5, ports are used to indirectly identify processes. They are a form of demux keys as de- fined in Section 1.3.1.) The port number is usually some well-known number specific to the service being offered; for example, Web servers commonly accept connections on port 80. The listen operation then defines how many connections can be pending on the specified socket. Finally, the accept operation carries out the passive open. It is a blocking operation that does not return until a remote participant has established a connection, and when it does complete, it returns a new socket that corresponds to this just-established connection, and the address argument contains the remote 1.4 Implementing Network Software 33 participant’s address. Note that when accept returns, the original socket that was given as an argument still exists and still corresponds to the passive open; it is used in future invocations of accept. On the client machine, the application process performs an active open; that is, it says who it wants to communicate with by invoking the following single operation: int connect(int socket, struct sockaddr *address, int addr len) This operation does not return until TCP has successfully established a connection, at which time the application is free to begin sending data. In this case, address contains the remote participant’s address. In practice, the client usually specifies only the remote participant’s address and lets the system fill in the local information. Whereas a server usually listens for messages on a well-known port, a client typically does not care which port it uses for itself; the OS simply selects an unused one. Once a connection is established, the application processes invoke the following two operations to send and receive data: int send(int socket, char *message, int msg len, int flags) int recv(int socket, char *buffer, int buf len, int flags) The first operation sends the given message over the specified socket, while the second operation receives a message from the specified socket into the given buffer. Both operations take a set of flags that control certain details of the operation. 1.4.2 Example Application We now show the implementation of a simple client/server program that uses the socket interface to send messages over a TCP connection. The program also uses other Unix networking utilities, which we introduce as we go. Our application allows a user on one machine to type in and send text to a user on another machine. It is a simplified version of the Unix talk program, which is similar to the program at the core of a Web chat room. Client We start with the client side, which takes the name of the remote machine as an argument. It calls the Unix utility gethostbyname to translate this name into the remote host’s IP address. The next step is to construct the address data structure (sin) expected by the socket interface. Notice that this data structure specifies that we’ll be using the socket to connect to the Internet (AF INET). In our example, we use TCP port 5432 as the well-known server port; this happens to be a port that has not been assigned to any other Internet service. The final step in setting up the connection is to call socket and connect. Once the connect operation returns, the connection is established and the 34 1 Foundation client program enters its main loop, which reads text from standard input and sends it over the socket. #include #include #include #include #include #define SERVER_PORT 5432 #define MAX_LINE 256 int main(int argc, char * argv[]) { FILE *fp; struct hostent *hp; struct sockaddr_in sin; char *host; char buf[MAX_LINE]; int s; int len; if (argc==2) { host = argv[1]; } else { fprintf(stderr, "usage: simplex-talk host\n"); exit(1); } /* translate host name into peer's IP address */ hp = gethostbyname(host); if (!hp) { fprintf(stderr, "simplex-talk: unknown host: %s\n", host); exit(1); } /* build address data structure */ bzero((char *)&sin, sizeof(sin)); sin.sin_family = AF_INET; bcopy(hp->h_addr, (char *)&sin.sin_addr, hp->h_length); sin.sin_port = htons(SERVER_PORT); /* active open */ if ((s = socket(PF_INET, SOCK_STREAM, 0)) < 0) { perror("simplex-talk: socket"); exit(1); 1.4 Implementing Network Software 35 } if (connect(s, (struct sockaddr *)&sin, sizeof(sin)) < 0) { perror("simplex-talk: connect"); close(s); exit(1); } /* main loop: get and send lines of text */ while (fgets(buf, sizeof(buf), stdin)) { buf[MAX_LINE-1] = '\0'; len = strlen(buf) + 1; send(s, buf, len, 0); } } Server The server is equally simple. It first constructs the address data structure by filling in its own port number (SERVER PORT). By not specifying an IP address, the application program is willing to accept connections on any of the local host’s IP addresses. Next, the server performs the preliminary steps involved in a passive open: creates the socket, binds it to the local address, and sets the maximum number of pending connections to be allowed. Finally, the main loop waits for a remote host to try to connect, and when one does, receives and prints out the characters that arrive on the connection. #include #include #include #include #include #define SERVER_PORT 5432 #define MAX_PENDING 5 #define MAX_LINE 256 int main() { struct sockaddr_in sin; char buf[MAX_LINE]; int len; int s, new_s; /* build address data structure */ bzero((char *)&sin, sizeof(sin)); sin.sin_family = AF_INET; sin.sin_addr.s_addr = INADDR_ANY; sin.sin_port = htons(SERVER_PORT); 36 1 Foundation /* setup passive open */ if ((s = socket(PF_INET, SOCK_STREAM, 0)) < 0) { perror("simplex-talk: socket"); exit(1); } if ((bind(s, (struct sockaddr *)&sin, sizeof(sin))) < 0) { perror("simplex-talk: bind"); exit(1); } listen(s, MAX_PENDING); /* wait for connection, then receive and print text */ while(1) { if ((new_s = accept(s, (struct sockaddr *)&sin, &len)) < 0){ perror("simplex-talk: accept"); exit(1); } while (len = recv(new_s, buf, sizeof(buf), 0)) fputs(buf, stdout); close(new_s); } } 1.4.3 Protocol Implementation Issues As mentioned at the beginning of this section, the way application programs interact with the underlying network is similar to the way a high-level protocol interacts with a low-level protocol. For example, TCP needs an interface to send outgoing messages to IP, and IP needs to be able to deliver incoming messages to TCP. This is exactly the service interface introduced in Section 1.3.1. Since we already have a network API (e.g., sockets), we might be tempted to use this same interface between every pair of protocols in the protocol stack. Although certainly an option, in practice the socket interface is not used in this way. The reason is that there are inefficiencies built into the socket interface that protocol implementers are not willing to tolerate. Application programmers tolerate them because they sim- plify their programming task and because the inefficiency only has to be tolerated once, but protocol implementers are often obsessed with performance and must worry about getting a message through several layers of protocols. The rest of this section discusses the two primary differences between the network API and the protocol-to-protocol interface found lower in the protocol graph. Process Model Most operating systems provide an abstraction called a process, or alternatively, a thread. Each process runs largely independently of other processes, and the OS is 1.4 Implementing Network Software 37 (a) (b) Figure 1.16 Alternative process models: (a) process-per-protocol; (b) process-per- message. responsible for making sure that resources, such as address space and CPU cycles, are allocated to all the current processes. The process abstraction makes it fairly straightforward to have a lot of things executing concurrently on one machine; for example, each user application might execute in its own process, and various things inside the OS might execute as other processes. When the OS stops one process from executing on the CPU and starts up another one, we call the change a context switch. When designing the network subsystem, one of the first questions to answer is, “Where are the processes?” There are essentially two choices, as illustrated in Figure 1.16. In the first, which we call the process-per-protocol model, each protocol is implemented by a separate process. This means that as a message moves up or down the protocol stack, it is passed from one process/protocol to another—the process that implements protocol i processes the message, then passes it to protocol i − 1, and so on. How one process/protocol passes a message to the next process/protocol depends on the support the host OS provides for interprocess communication. Typically, there is a simple mechanism for enqueuing a message with a process. The important point, however, is that a context switch is required at each level of the protocol graph— typically a time-consuming operation. The alternative, which we call the process-per-message model, treats each pro- tocol as a static piece of code and associates the processes with the messages. 38 1 Foundation That is, when a message arrives from the network, the OS dispatches a process that it makes responsible for the message as it moves up the protocol graph. At each level, the procedure that implements that protocol is invoked, which eventually results in the procedure for the next protocol being invoked, and so on. For outbound messages, the application’s process invokes the necessary procedure calls until the message is deliv- ered. In both directions, the protocol graph is traversed in a sequence of procedure calls. Although the process-per-protocol model is sometimes easier to think about— I implement my protocol in my process, and you implement your protocol in your process—the process-per-message model is generally more efficient for a simple reason: A procedure call is an order of magnitude more efficient than a context switch on most computers. The former model requires the expense of a context switch at each level, while the latter model costs only a procedure call per level. Now think about the relationship between the service interface as defined above and the process model. For an outgoing message, the high-level protocol invokes a send operation on the low-level protocol. Because the high-level protocol has the message in hand when it calls send, this operation can be easily implemented as a procedure call; no context switch is required. For incoming messages, however, the high-level protocol invokes the receive operation on the low-level protocol, and then must wait for a message to arrive at some unknown future time; this basically forces a context switch. In other words, the process running in the high-level protocol receives a message from the process running in the low-level protocol. This isn’t a big deal if only the application process receives messages from the network subsystem—in fact, it’s the right interface for the network API since application programs already have a process- centric view of the world—but it does have a significant impact on performance if such a context switch occurs at each layer of the protocol stack. It is for this reason that most protocol implementations replace the receive oper- ation with a deliver operation. That is, the low-level protocol does an upcall—a pro- cedure call up the protocol stack—to deliver the message to the high-level protocol. Figure 1.17 shows the resulting interface between two adjacent protocols, TCP and IP in this case. In general, messages move down the protocol graph through a sequence of send operations, and up the protocol graph through a sequence of deliver operations. Message Buffers A second inefficiency of the socket interface is that the application process provides the buffer that contains the outbound message when calling send, and similarly it provides the buffer into which an incoming message is copied when invoking the receive operation. This forces the topmost protocol to copy the message from the ap- plication’s buffer into a network buffer, and vice versa, as shown in Figure 1.18. It turns out that copying data from one buffer to another is one of the most expensive things a 1.4 Implementing Network Software 39 sendIP (message) deliverTCP (message) TCP IP Figure 1.17 Protocol-to-protocol interface. send() deliver() Topmost protocol Application process Figure 1.18 Copying incoming/outgoing messages between application buffer and network buffer. protocol implementation can do. This is because while processors are becoming faster at an incredible pace, memory is not getting faster as quickly as processors are. Instead of copying message data from one buffer to another at each layer in the protocol stack, most network subsystems define an abstract data type for messages that is shared by all protocols in the protocol graph. Not only does this abstraction permit messages to be passed up and down the protocol graph without copying, but it usually provides copy-free ways of manipulating messages in other ways, such as adding and stripping headers, fragmenting large messages into a set of small messages, and reassembling a collection of small messages into a single large message. The exact form of this message abstraction differs from OS to OS, but it generally involves a linked list of pointers to message buffers, similar to the one shown in Figure 1.19. We leave it as an exercise for you to define a general copy-free message abstraction. 40 1 Foundation Figure 1.19 Example message data structure. 1.5 Performance Up to this point, we have focused primar- ily on the functional aspects of networks. Like any computer system, however, com- puter networks are also expected to perform well, since the effectiveness of computations distributed over the network often depends directly on the efficiency with which the net- work delivers the computation’s data. While the old programming adage “First get it right and then make it fast” is valid in many settings, in networking it is usually neces- sary to “design for performance.” It is there- fore important to understand the various factors that impact network performance. 1.5.1 Bandwidth and Latency Network performance is measured in two fundamental ways: bandwidth (also called throughput) and latency (also called delay). The bandwidth of a network is given by the number of bits that can be transmitted over the network in a certain period of time. For example, a network might have a band- width of 10 million bits/second (Mbps), meaning that it is able to deliver 10 million Bandwidth and Throughput Bandwidth and throughput are two of the most confusing terms used in networking. While we could try to give you a precise definition of each term, it is important that you know how other people might use them and for you to be aware that they are often used interchange- ably. First of all, bandwidth is lit- erally a measure of the width of a frequency band. For example, a voice-grade telephone line sup- ports a frequency band ranging from 300 to 3300 Hz; it is said to have a bandwidth of 3300 Hz −300 Hz = 3000 Hz. If you see the word “bandwidth” used in a situa- tion in which it is being measured in hertz, then it probably refers to the range of signals that can be accommodated. When we talk about the band- width of a communication link, we 1.5 Performance 41 bits every second. It is sometimes useful to think of bandwidth in terms of how long it takes to transmit each bit of data. On a 10-Mbps network, for example, it takes 0.1 microsecond (μs) to transmit each bit. While you can talk about the bandwidth of the network as a whole, sometimes you want to be more precise, focusing, for example, on the bandwidth of a single physical link or of a logical process-to-process channel. At the physical level, band- width is constantly improving, with no end in sight. Intuitively, if you think of a second of time as a distance you could measure with a ruler, and bandwidth as how many bits fit in that distance, then you can think of each bit as a pulse of some width. For example, each bit on a 1-Mbps link is 1 μs wide, while each bit on a 2-Mbps link is 0.5 μs wide, as illustrated in Figure 1.20. The more sophisticated the transmitting and receiving technology, the narrower each bit can become, and thus, the higher the normally refer to the number of bits per second that can be trans- mitted on the link. We might say that the bandwidth of an Ethernet is 10 Mbps. A useful distinction might be made, however, between the bandwidth that is available on the link and the number of bits per second that we can actually trans- mit over the link in practice. We tend to use the word “throughput” to refer to the measured perfor- mance of a system. Thus, because of various inefficiencies of implemen- tation, a pair of nodes connected by a link with a bandwidth of 10 Mbps might achieve a throughput of only 2 Mbps. This would mean that an application on one host could send data to the other host at 2 Mbps. Finally, we often talk about the bandwidth requirements of an application—the number of bits bandwidth. For logical process-to-process channels, bandwidth is also influenced by other factors, including how many times the software that implements the channel has to handle, and possibly transform, each bit of data. The second performance metric, la- tency, corresponds to how long it takes a message to travel from one end of a network to the other. (As with bandwidth, we could be focused on the latency of a single link or an end-to-end channel.) Latency is mea- sured strictly in terms of time. For example, a transcontinental network might have a la- tency of 24 milliseconds (ms); that is, it takes a message 24 ms to travel from one end of North America to the other. There are many situations in which it is more important to know how long it takes to send a message from one end of a network to the other and back, rather than the one-way latency. We call this the round-trip time (RTT) of the network. We often think of latency as having three components. First, there is the speed- of-light propagation delay. This delay oc- curs because nothing, including a bit on a wire, can travel faster than the speed 42 1 Foundation 1 second (a) 1 second (b) Figure 1.20 Bits transmitted at a particular bandwidth can be regarded as having some width: (a) bits transmitted at 1 Mbps (each bit 1 µs wide); (b) bits transmitted at 2 Mbps (each bit 0.5 µs wide). of light. If you know the distance between two points, you can calculate the speed-of- light latency, although you have to be careful because light travels across different medi- ums at different speeds: It travels at 3.0×108 m/s in a vacuum, 2.3×108 m/s in a cable, and 2.0×108 m/s in a fiber.Second, there is the amount of time it takes to transmit a unit of data. This is a function of the network bandwidth and the size of the packet in which the data is carried. Third, there may be queuing delays inside the network, since packet switches generally need to store packets for some time before forwarding them on an outbound link, as discussed in Section 1.2.2. So, we could define the total latency as Latency = Propagation + Transmit + Queue Propagation = Distance/SpeedOfLight Transmit = Size/Bandwidth where Distance is the length of the wire over which the data will travel, SpeedOfLight is the effective speed of light over that wire, Size is the size of the packet, and Bandwidth is the bandwidth at which the packet is transmitted. Note that if the message contains only one bit and we are talking about a single link (as opposed to a whole network), then the Transmit and Queue terms are not relevant, and latency corresponds to the propagation delay only. Bandwidth and latency combine to de- fine the performance characteristics of a given link or channel. Their relative impor- tance, however, depends on the application. For some applications, latency dominates per second that it needs to trans- mit over the network to perform acceptably. For some applications, this might be “whatever I can get”; for others, it might be some fixed number (preferably no more than the available link bandwidth); and for others, it might be a number that varies with time. We will pro- vide more on this topic later in this section. 1.5 Performance 43 bandwidth. For example, a client that sends a 1-byte message to a server and receives a 1-byte message in return is latency bound. Assuming that no serious computation is involved in preparing the response, the application will perform much differently on a transcontinental channel with a 100-ms RTT than it will on an across-the-room channel with a 1-ms RTT. Whether the channel is 1 Mbps or 100 Mbps is relatively in- significant, however, since the former implies that the time to transmit a byte (Transmit) is 8 μs and the latter implies Transmit = 0.08 μs. In contrast, consider a digital library program that is being asked to fetch a 25-megabyte (MB) image—the more bandwidth that is available, the faster it will be able to return the image to the user. Here, the bandwidth of the channel dominates performance. To see this, suppose that the channel has a bandwidth of 10 Mbps. It will take 20 seconds to transmit the image, making it relatively unimportant if the image is on the other side of a 1-ms channel or a 100-ms channel; the difference between a 20.001-second response time and a 20.1-second response time is negligible. Figure 1.21 gives you a sense of how latency or bandwidth can dominate perfor- mance in different circumstances. The graph shows how long it takes to move objects 10,000 5000 2000 1000 500 200 100 50 20 10 5 2 1 10010 RTT (ms) 1-MB object, 1.5-Mbps link 1-MB object, 10-Mbps link 2-KB object, 1.5-Mbps link 2-KB object, 10-Mbps link 1-byte object, 1.5-Mbps link 1-byte object, 10-Mbps linkPerceived latency (ms) Figure 1.21 Perceived latency (response time) versus round-trip time for various ob- ject sizes and link speeds. 44 1 Foundation of various sizes (1 byte, 2 KB, 1 MB) across networks with RTTs ranging from 1 to 100 ms and link speeds of either 1.5 or 10 Mbps. We use logarithmic scales to show relative performance. For a 1-byte object (say, a keystroke), latency remains almost exactly equal to the RTT, so that you cannot distinguish between a 1.5-Mbps network and a 10-Mbps network. For a 2-KB object (say, an email message), the link speed makes quite a difference on a 1-ms RTT network but a negligible difference on a 100- ms RTT network. And for a 1-MB object (say, a digital image), the RTT makes no difference—it is the link speed that dominates performance across the full range of RTT. Note that throughout this book we use the terms latency and delay in a generic way, that is, to denote how long it takes to perform a particular function such as delivering a message or moving an object. When we are referring to the specific amount of time it takes a signal to propagate from one end of a link to another, we use the term propagation delay. Also, we make it clear in the context of the discussion whether we are referring to the one-way latency or the round-trip time. As an aside, computers are becoming so fast that when we connect them to net- works, it is sometimes useful to think, at least figuratively, in terms of instructions per mile. Consider what happens when a com- puter that is able to execute 1 billion instruc- tions per second sends a message out on a channel with a 100-ms RTT. (To make the math easier, assume that the message covers a distance of 5000 miles.) If that computer sits idle the full 100 ms waiting for a reply message, then it has forfeited the ability to execute 100 million instructions, or 20,000 instructions per mile. It had better have been worth going over the network to justify this waste. 1.5.2 Delay × Bandwidth Product It is also useful to talk about the product of these two metrics, often called the delay ×bandwidth product. Intuitively, if we think of a channel between a pair of processes as a hollow pipe (see Figure 1.22), where How Big Is a Mega? There are several pitfalls you need to be aware of when working with the common units of networking— MB, Mbps, KB, and Kbps. The first is to distinguish carefully be- tween bits and bytes. Throughout this book, we always use a lower- case b for bits and a capital B for bytes. The second is to be sure you are using the appropriate definition of mega (M) and kilo (K). Mega, for example, can mean either 220 or 106. Similarly, kilo can be either 210 or 103. What is worse, in net- working we typically use both def- initions. Here’s why. Network bandwidth, which is often specified in terms of Mbps, is typically governed by the speed of the clock that paces the trans- mission of the bits. A clock that is 1.5 Performance 45 Bandwidth Delay Figure 1.22 Network as a pipe. the latency corresponds to the length of the pipe and the bandwidth gives the diameter of the pipe, then the delay × bandwidth product gives the volume of the pipe—the number of bits it holds. Said another way, if latency (measured in time) corresponds to the length of the pipe, then given the width of each bit (also measured in time), you can calculate how many bits fit in the pipe. For example, a transcontinental chan- nel with a one-way latency of 50 ms and a bandwidth of 45 Mbps is able to hold running at 10 MHz is used to trans- mit bits at 10 Mbps. Because the mega in MHz means 106 hertz, Mbps is usually also defined as 106 bits per second. (Similarly, Kbps is 103 bits per second.) On the other hand, when we talk about a mes- sage that we want to transmit, we often give its size in kilobytes. Be- cause messages are stored in the computer’s memory, and memory is typically measured in powers of two, the K in KB is usually taken to mean 210. (Similarly, MB usu- ally means 220.) When you put the two together, it is not un- common to talk about sending a 32-KB message over a 10-Mbps channel, which should be inter- preted to mean 32 × 210 × 8 bits are being transmitted at a rate of 50 × 10−3 seconds × 45 × 106 bits/second = 2.25 × 106 bits or approximately 280 KB of data. In other words, this example channel (pipe) holds as many bytes as the memory of a personal computer from the early 1980s could hold. The delay × bandwidth product is im- portant to know when constructing high- performance networks because it corre- sponds to how many bits the sender must transmit before the first bit arrives at the receiver. If the sender is expecting the re- ceiver to somehow signal that bits are start- ing to arrive, and it takes another channel latency for this signal to propagate back to the sender (i.e., we are interested in the channel’s RTT rather than just its one-way latency), then the sender can send up to two delay × bandwidth’s worth of data be- fore hearing from the receiver that all is well. The bits in the pipe are said to be “in flight,” which means that if the receiver tells the sender to stop transmitting, it might re- ceive up to a delay × bandwidth’s worth of 46 1 Foundation data before the sender manages to respond. In our example above, that amount corre- sponds to 5.5 × 106 bits (671 KB) of data. On the other hand, if the sender does not fill the pipe—send a whole delay × bandwidth product’s worth of data before it stops to wait for a signal—the sender will not fully utilize the network. Note that most of the time we are interested in the RTT scenario, which we simply refer to as the delay × bandwidth product, without explicitly saying that this product is multiplied by two. Again, whether the “delay” in “delay × band- width” means one-way latency or RTT is made clear by the context. 1.5.3 High-Speed Networks The bandwidths available on today’s net- works are increasing at a dramatic rate, and there is eternal optimism that network band- width will continue to improve. This causes network designers to start thinking about what happens in the limit, or stated another way, what is the impact on network design of having infinite bandwidth available. Although high-speed networks bring a dramatic change in the bandwidth avail- able to applications, in many respects their impact on how we think about networking comes in what does not change as band- width increases: the speed of light. To quote Scotty from Star Trek, “You cannae change the laws of physics.” In other words, “high speed” does not mean that latency improves at the same rate as bandwidth; the transcon- tinental RTT of a 1-Gbps link is the same 100 ms as it is for a 1-Mbps link. 10×106 bits per second. This is the interpretation we use throughout the book, unless explicitly stated otherwise. The good news is that many times we are satisfied with a back-of-the-envelope calculation, in which case it is perfectly reason- able to pretend that a byte has 10 bits in it (making it easy to convert between bits and bytes) and that 106 is really equal to 220 (making it easy to convert between the two definitions of mega). Notice that the first approximation introduces a 20% error, while the latter intro- duces only a 5% error. To help you in your quick- and-dirty calculations, 100 ms is a reasonable number to use for a cross-country round-trip time—at least when the country in question is the United States—and 1 ms is a good approximation of an RTT across a local area network. In the case of the former, we increase the 48-ms round-trip time implied by the speed of light over a fiber to 100 ms because there are, as we have said, other sources of delay, such as the processing time in the switches inside the network. You can also be sure that the path taken by the fiber between two points will not be a straight line. 1.5 Performance 47 1-Mbps cross-country link 1-Gbps cross-country link (b) 1 MB . .. (a) Source Source .1 Mb .1 Mb .1 Mb .1 Mb Destination Destination Figure 1.23 Relationship between bandwidth and latency. With a 1-MB file, (a) the 1-Mbps link has 80 pipes full of data; (b) the 1-Gbps link has 1/12 of one pipe full of data. To appreciate the significance of ever-increasing bandwidth in the face of fixed latency, consider what is required to transmit a 1-MB file over a 1-Mbps network versus over a 1-Gbps network, both of which have an RTT of 100 ms. In the case of the 1-Mbps network, it takes 80 round-trip times to transmit the file; during each RTT, 1.25% of the file is sent. In contrast, the same 1-MB file doesn’t even come close to filling 1 RTT’s worth of the 1-Gbps link, which has a delay × bandwidth product of 12.5 MB. Figure 1.23 illustrates the difference between the two networks. In effect, the 1-MB file looks like a stream of data that needs to be transmitted across a 1-Mbps network, while it looks like a single packet on a 1-Gbps network. To help drive this point home, consider that a 1-MB file is to a 1-Gbps network what a 1-KB packet is to a 1-Mbps network.◮ Another way to think about the situation is that more data can be transmitted during each RTT on a high-speed network, so much so that a single RTT becomes a significant amount of time. Thus, while you wouldn’t think twice about the difference 48 1 Foundation between a file transfer taking 101 RTTs rather than 100 RTTs (a relative difference of only 1%), suddenly the difference between 1 RTT and 2 RTTs is significant—a 100% increase. In other words, latency, rather than throughput, starts to dominate our thinking about network design. Perhaps the best way to understand the relationship between throughput and latency is to return to basics. The effective end-to-end throughput that can be achieved over a network is given by the simple relationship Throughput = TransferSize/TransferTime where TransferTime includes not only the elements of one-way Latency identified earlier in this section, but also any additional time spent requesting or setting up the transfer. Generally, we represent this relationship as TransferTime = RTT + 1/Bandwidth × TransferSize We use RTT in this calculation to account for a request message being sent across the network and the data being sent back. For example, consider a situation where a user wants to fetch a 1-MB file across a 1-Gbps network with a round-trip time of 100 ms. The TransferTime includes both the transmit time for 1 MB (1/1 Gbps × 1MB= 8 ms), and the 100-ms RTT, for a total transfer time of 108 ms. This means that the effective throughput will be 1MB/108 ms = 74.1 Mbps not 1 Gbps. Clearly, transferring a larger amount of data will help improve the ef- fective throughput, where in the limit, an infinitely large transfer size will cause the effective throughput to approach the network bandwidth. On the other hand, having to endure more than 1 RTT—for example, to retransmit missing packets—will hurt the effective throughput for any transfer of finite size and will be most noticeable for small transfers. 1.5.4 Application Performance Needs The discussion in this section has taken a network-centric view of performance; that is, we have talked in terms of what a given link or channel will support. The unstated assumption has been that application programs have simple needs—they want as much bandwidth as the network can provide. This is certainly true of the aforementioned digital library program that is retrieving a 25-MB image; the more bandwidth that is available, the faster the program will be able to return the image to the user. However, some applications are able to state an upper limit on how much band- width they need. Video applications are a prime example. Suppose you want to stream 1.5 Performance 49 a video image that is one-quarter the size of a standard TV image; that is, it has a resolution of 352 by 240 pixels. If each pixel is represented by 24 bits of information, as would be the case for 24-bit color, then the size of each frame would be (352 × 240 × 24)/8 = 247.5KB If the application needs to support a frame rate of 30 frames per second, then it might request a throughput rate of 75 Mbps. The ability of the network to provide more bandwidth is of no interest to such an application because it has only so much data to transmit in a given period of time. Unfortunately, the situation is not as simple as this example suggests. Because the difference between any two adjacent frames in a video stream is often small, it is possible to compress the video by transmitting only the differences between adjacent frames. This compressed video does not flow at a constant rate, but varies with time according to factors such as the amount of action and detail in the picture and the compression algorithm being used. Therefore, it is possible to say what the average bandwidth requirement will be, but the instantaneous rate may be more or less. The key issue is the time interval over which the average is computed. Suppose that this example video application can be compressed down to the point that it needs only 2 Mbps, on average. If it transmits 1 megabit in a 1-second interval and 3 megabits in the following 1-second interval, then over the 2-second interval it is transmitting at an average rate of 2 Mbps; however, this will be of little consolation to a channel that was engineered to support no more than 2 megabits in any one second. Clearly, just knowing the average bandwidth needs of an application will not always suffice. Generally, however, it is possible to put an upper bound on how big of a burst an application like this is likely to transmit. A burst might be described by some peak rate that is maintained for some period of time. Alternatively, it could be described as the number of bytes that can be sent at the peak rate before reverting to the average rate or some lower rate. If this peak rate is higher than the available channel capacity, then the excess data will have to be buffered somewhere, to be transmitted later. Knowing how big of a burst might be sent allows the network designer to allocate sufficient buffer capacity to hold the burst. We will return to the subject of describing bursty traffic accurately in Chapter 6. Analogous to the way an application’s bandwidth needs can be something other than “all it can get,” an application’s delay requirements may be more complex than simply “as little delay as possible.” In the case of delay, it sometimes doesn’t matter so much whether the one-way latency of the network is 100 ms or 500 ms as how much the latency varies from packet to packet. The variation in latency is called jitter. Consider the situation in which the source sends a packet once every 33 ms, as would be the case for a video application transmitting frames 30 times a second. 50 1 Foundation Network { Interpacket gap Packet source Packet sink 1234 1234 Figure 1.24 Network-induced jitter. If the packets arrive at the destination spaced out exactly 33 ms apart, then we can deduce that the delay experienced by each packet in the network was exactly the same. If the spacing between when packets arrive at the destination—sometimes called the interpacket gap—is variable, however, then the delay experienced by the sequence of packets must have also been variable, and the network is said to have introduced jitter into the packet stream, as shown in Figure 1.24. Such variation is generally not introduced in a single physical link, but it can happen when packets experience different queuing delays in a multihop packet-switched network. This queuing delay corresponds to the Queue component of latency defined earlier in this section, which varies with time. To understand the relevance of jitter, suppose that the packets being transmitted over the network contain video frames, and in order to display these frames on the screen the receiver needs to receive a new one every 33 ms. If a frame arrives early, then it can simply be saved by the receiver until it is time to display it. Unfortunately, if a frame arrives late, then the receiver will not have the frame it needs in time to update the screen, and the video quality will suffer; it will not be smooth. Note that it is not necessary to eliminate jitter, only to know how bad it is. The reason for this is that if the receiver knows the upper and lower bounds on the latency that a packet can experience, it can delay the time at which it starts playing back the video (i.e., displays the first frame) long enough to ensure that in the future it will always have a frame to display when it needs it. The receiver delays the frame, effectively smoothing out the jitter, by storing it in a buffer. We return to the topic of jitter in Chapter 9. 1.6 Summary Computer networks like the Internet have experienced enormous growth over the past decade and are now positioned to provide a wide range of services—remote file access, digital libraries, videoconferencing—to hundreds of millions of users. Much of this growth can be attributed to the general-purpose nature of computer networks, and in particular to the ability to add new functionality to the network by writing Open Issue: Ubiquitous Networking 51 software that runs on affordable, high-performance computers. With this in mind, the overriding goal of this book is to describe computer networks in such a way that when you finish reading it, you should feel that if you had an army of programmers at your disposal, you could actually build a fully functional computer network from the ground up. This chapter lays the foundation for realizing this goal. The first step we have taken toward this goal is to carefully identify exactly what we expect from a network. For example, a network must first provide cost- effective connectivity among a set of computers. This is accomplished through a nested interconnection of nodes and links, and by sharing this hardware base through the use of statistical multiplexing. This results in a packet-switched network, on top of which we then define a collection of process-to-process communication services. The second step is to define a layered architecture that will serve as a blueprint for our design. The central objects of this architecture are network protocols. Protocols both provide a communication service to higher-level protocols and define the form and meaning of messages exchanged with their peers running on other machines. We have briefly surveyed two of the most widely used architectures: the OSI architecture and the Internet architecture. This book most closely follows the Internet architecture, both in its organization and as a source of examples. The third step is to implement the network’s protocols and application programs, usually in software. Both protocols and applications need an interface by which they invoke the services of other protocols in the network subsystem. The socket interface is the most widely used interface between application programs and the network subsys- tem, but a slightly different interface is typically used within the network subsystem. Finally, the network as a whole must offer high performance, where the two performance metrics we are most interested in are latency and throughput. As we will see in later chapters, it is the product of these two metrics—the so-called delay ×bandwidth product—that often plays a critical role in protocol design. There is little doubt that computer networks are becoming an integral part of the everyday lives of vast numbers of people. What began over 20 years ago as experimental sys- tems like the ARPANET—connecting mainframe computers over long- OPEN ISSUE Ubiquitous Networking distance telephone lines—has turned into big business. And where there is big busi- ness, there are lots of players. In this case, there is the computing industry, which has 52 1 Foundation become increasingly involved in supporting packet-switched networking products; the telephone carriers, which recognize the market for carrying all sorts of data, not just voice; and the cable TV industry, which currently owns the entertainment portion of the market. Assuming that the goal is ubiquitous networking—to bring the network into every household—the first problem that must be addressed is how to establish the necessary physical links. Although it could be argued that the ultimate answer is to bring an optical fiber into every home, at an estimated $1000 per house and 100 million homes in the U.S. alone, this is a $100 billion proposition. The most widely discussed alternatives make use of either the existing cable TV facilities or the copper pairs used to deliver telephone service. Each of these approaches has its own set of problems. For example, today’s cable facilities are asymmetric—you can deliver 150 channels into every home, but the outgoing bandwidth is severely limited. Such asymmetry implies that there are a small number of information providers, but that most of us are simply information consumers. Many people would argue that in a democracy we should all have an equal opportunity to provide information. Digital subscriber line (DSL) technology need not be asymmetric, but can only offer high-bandwidth connections to a subset of consumers over the existing telephone wires. How the struggle between the computer companies, the telephone companies, and the cable industry will play out in the marketplace is anyone’s guess. (If we knew the answer, we’d be charging a lot more for this book.) All we know is that there are many technical obstacles—issues of connectivity, levels of service, performance, reliability, and fairness—that stand between the current state of the art and the sort of global, ubiquitous, heterogeneous network that we believe is possible and desirable. It is these challenges that are the focus of this book. FURTHER READING Computer networks are not the first communication-oriented technology to have found their way into the everyday fabric of our society. For example, the early part of this century saw the introduction of the telephone, and then during the 1950s tele- vision became widespread. When considering the future of networking—how widely it will spread and how we will use it—it is instructive to study this history. Our first reference is a good starting point for doing this (the entire issue is devoted to the first 100 years of telecommunications). The second and third papers are the seminal papers on the OSI and Internet architectures, respectively. The Zimmerman paper introduces the OSI architecture, and the Clark paper is a retrospective. The final two papers are not specific to networking, Further Reading 53 but ones that every systems person should read. The Saltzer et al. paper motivates and describes one of the most widely applied rules of system design—the end-to-end argument. The paper by Mashey describes the thinking behind RISC architectures; as we will soon discover, making good judgments about where to place functionality in a complex system is what system design is all about. ■ Pierce, J. Telephony—a personal view. IEEE Communications 22(5):116–120, May 1984. ■ Zimmerman, H. OSI reference model—the ISO model of architecture for open systems interconnection. IEEE Transactions on Communications COM- 28(4):425–432, April 1980. ■ Clark, D. The design philosophy of the DARPA Internet protocols. Proceed- ings of the SIGCOMM ’88 Symposium, pages 106–114, August 1988. ■ Saltzer, J., D. Reed, and D. Clark. End-to-end arguments in system design. ACM Transactions on Computer Systems 2(4):277–288, November 1984. ■ Mashey, J. RISC, MIPS, and the motion of complexity. UniForum 1986 Con- ference Proceedings, pages 116–124, 1986. Several texts offer an introduction to computer networking: Stallings gives an encyclopedic treatment of the subject, with an emphasis on the lower levels of the OSI hierarchy [Sta00a]; Tanenbaum uses the OSI architecture as an organizational model [Tan02]; Comer gives an overview of the Internet architecture [Com00]; and Bertsekas and Gallager discuss networking from a performance modeling perspective [BG92]. To put computer networking into a larger context, two books—one dealing with the past and the other looking toward the future—are must reading. The first is Holzmann and Pehrson’s The Early History of Data Networks [HP95]. Surprisingly, many of the ideas covered in the book you are now reading were invented during the 1700s. The second is Realizing the Information Future: The Internet and Beyond, a book prepared by the Computer Science and Telecommunications Board of the National Research Council [NRC94]. To follow the history of the Internet from its beginning, you are encouraged to peruse the Internet’s Request for Comments (RFC) series of documents. These docu- ments, which include everything from the TCP specification to April Fools’ jokes, are retrievable at http://www.ietf.org/rfc.html. For example, the protocol specifications for TCP, UDP, and IP are available in RFC 793, 768, and 791, respectively. To gain a better appreciation for the Internet philosophy and culture, two ref- erences are must reading; both are also quite entertaining. Padlipsky gives a good description of the early days, including a pointed comparison of the Internet and OSI 54 1 Foundation architectures [Pad85]. For a more up-to-date account of what really happens behind the scenes at the Internet Engineering Task Force, we recommend Boorsook’s article [Boo95]. There are a wealth of articles discussing various aspects of protocol implemen- tations. A good starting point is to understand two complete protocol implementation environments: the Stream mechanism from System V Unix [Rit84] and the x-kernel [HP91]. In addition, [LMKQ89] and [SW95] describe the widely used Berkeley Unix implementation of TCP/IP. More generally, there is a large body of work addressing the issue of structuring and optimizing protocol implementations. Clark was one of the first to discuss the relationship between modular design and protocol performance [Cla82]. Later pa- pers then introduce the use of upcalls in structuring protocol code [Cla85] and study the processing overheads in TCP [CJRS89]. Finally, [WM87] describes how to gain efficiency through appropriate design and implementation choices. Several papers have introduced specific techniques and mechanisms that can be used to improve protocol performance. For example, [HMPT89] describes some of the mechanisms used in the x-kernel, [MD93] discusses various implementations of demultiplexing tables, [VL87] introduces the timing wheel mechanism used to manage protocol events, and [DP93] describes an efficient buffer management strategy. Also, the performance of protocols running on parallel processors—locking is a key issue in such environments—is discussed in [BG93] and [NYKT94]. Because many aspects of protocol implementation depend on an understanding of the basics of operating systems, we recommend Finkel [Fin88], Bic and Shaw [BS88], and Tanenbaum [Tan01] for an introduction to OS concepts. Finally, we conclude the “Further Reading” section of each chapter with a set of live references, that is, URLs for locations on the World Wide Web where you can learn more about the topics discussed in that chapter. Since these references are live, it is possible that they will not remain active for an indefinite period of time. For this reason, we limit the set of live references at the end of each chapter to sites that either export software, provide a service, or report on the activities of an ongoing working group or standardization body. In other words, we only give URLs for the kinds of material that cannot easily be referenced using standard citations. For this chapter, we include four live references: ■ http://www.mkp.com: information about this book, including supplements, addendums, and so on ■ http://www.acm.org/sigcomm/sos.html: status of various networking stan- dards, including those of the IETF, ISO, and IEEE Exercises 55 ■ http://www.ietf.org/: information about the IETF and its working groups ■ http://www.cs.columbia.edu/˜ hgs/netbib/: searchable bibliography of network- related research papers EXERCISES 1 Use anonymous FTP to connect to ftp.isi.edu (directory in-notes), and retrieve the RFC index. Also retrieve the protocol specifications for TCP, IP, and UDP. 2 Look up the Web site http://www.cs.princeton.edu/nsg Here you can read about current network research under way at Princeton Uni- versity and see a picture of author Larry Peterson. Follow links to find a picture of author Bruce Davie. 3 Use a Web search tool to locate useful, general, and noncommercial information about the following topics: MBone, ATM, MPEG, IPv6, and Ethernet. 4 The Unix utility whois can be used to find the domain name corresponding to an organization, or vice versa. Read the man page documentation for whois and experiment with it. Try whois princeton.edu and whois princeton, for starters. 5 Calculate the total time required to transfer a 1000-KB file in the following cases, assuming an RTT of 100 ms, a packet size of 1 KB and an initial 2 × RTT of “handshaking” before data is sent. (a) The bandwidth is 1.5 Mbps, and data packets can be sent continuously. (b) The bandwidth is 1.5 Mbps, but after we finish sending each data packet we must wait one RTT before sending the next. (c) The bandwidth is “infinite,” meaning that we take transmit time to be zero, and up to 20 packets can be sent per RTT. (d) The bandwidth is infinite, and during the first RTT we can send one packet (21−1), during the second RTT we can send two packets (22−1), during the third we can send four (23−1), and so on. (A justification for such an exponential increase will be given in Chapter 6.) 56 1 Foundation 6 Calculate the total time required to transfer a 1.5-MB file in the following cases, assuming an RTT of 80 ms, a packet size of 1 KB and an initial 2 × RTT of “hand- shaking” before data is sent. (a) The bandwidth is 10 Mbps, and data packets can be sent continuously. (b) The bandwidth is 10 Mbps, but after we finish sending each data packet we must wait one RTT before sending the next. (c) The link allows infinitely fast transmit, but limits bandwidth such that only 20 packets can be sent per RTT. (d) Zero transmit time as in (c), but during the first RTT we can send one packet, during the second RTT we can send two packets, during the third we can send four = 23−1, and so on. (A justification for such an exponential increase will be given in Chapter 6.) 7 Consider a point-to-point link 2 km in length. At what bandwidth would propa- gation delay (at a speed of 2×108 m/s) equal transmit delay for 100-byte packets? What about 512-byte packets? 8 Consider a point-to-point link 50 km in length. At what bandwidth would propa- gation delay (at a speed of 2×108 m/s) equal transmit delay for 100-byte packets? What about 512-byte packets? 9 What properties of postal addresses would be likely to be shared by a network addressing scheme? What differences might you expect to find? What properties of telephone numbering might be shared by a network addressing scheme? 10 One property of addresses is that they are unique; if two nodes had the same address it would be impossible to distinguish between them. What other properties might be useful for network addresses to have? Can you think of any situations in which network (or postal or telephone) addresses might not be unique? 11 Give an example of a situation in which multicast addresses might be beneficial. 12 What differences in traffic patterns account for the fact that STDM is a cost- effective form of multiplexing for a voice telephone network and FDM is a cost- effective form of multiplexing for television and radio networks, yet we reject both as not being cost-effective for a general-purpose computer network? 13 How “wide” is a bit on a 1-Gbps link? How long is a bit in copper wire, where the speed of propagation is 2.3 × 108 m/s? Exercises 57 14 How long does it take to transmit x KB over a y-Mbps link? Give your answer as a ratio of x and y. 15 Suppose a 100-Mbps point-to-point link is being set up between Earth and a new lunar colony. The distance from the moon to Earth is approximately 385,000 km, and data travels over the link at the speed of light—3 × 108 m/s. (a) Calculate the minimum RTT for the link. (b) Using the RTT as the delay, calculate the delay × bandwidth product for the link. (c) What is the significance of the delay × bandwidth product computed in (b)? (d) A camera on the lunar base takes pictures of Earth and saves them in digital format to disk. Suppose Mission Control on Earth wishes to download the most current image, which is 25 MB. What is the minimum amount of time that will elapse between when the request for the data goes out and the transfer is finished? 16 Suppose a 128-Kbps point-to-point link is set up between Earth and a rover on Mars. The distance from Earth to Mars (when they are closest together) is approx- imately 55 Gm, and data travels over the link at the speed of light—3 × 108 m/s. (a) Calculate the minimum RTT for the link. (b) Calculate the delay × bandwidth product for the link. (c) A camera on the rover takes pictures of its surroundings and sends these to Earth. How quickly after a picture is taken can it reach Mission Control on Earth? Assume that each image is 5 Mb in size. 17 For each of the following operations on a remote file server, discuss whether they are more likely to be delay sensitive or bandwidth sensitive. (a) Open a file. (b) Read the contents of a file. (c) List the contents of a directory. (d) Display the attributes of a file. 18 Calculate the latency (from first bit sent to last bit received) for the following: (a) 10-Mbps Ethernet with a single store-and-forward switch in the path, and a packet size of 5000 bits. Assume that each link introduces a propagation delay of 10 μs and that the switch begins retransmitting immediately after it has finished receiving the packet. 58 1 Foundation (b) Same as (a) but with three switches. (c) Same as (a) but assume the switch implements “cut-through” switching: It is able to begin retransmitting the packet after the first 200 bits have been received. 19 Calculate the latency (from first bit sent to last bit received) for the following: (a) 1-Gbps Ethernet with a single store-and-forward switch in the path, and a packet size of 5000 bits. Assume that each link introduces a propagation delay of 10 μs and that the switch begins retransmitting immediately after it has finished receiving the packet. (b) Same as (a) but with three switches. (c) Same as (b) but assume the switch implements “cut-through” switching: It is able to begin retransmitting the packet after the first 128 bits have been received. 20 Calculate the effective bandwidth for the following cases. For (a) and (b) assume there is a steady supply of data to send; for (c) simply calculate the average over 12 hours. (a) 10-Mbps Ethernet through three store-and-forward switches as in Exercise 18(b). Switches can send on one link while receiving on the other. (b) Same as (a) but with the sender having to wait for a 50-byte acknowledgment packet after sending each 5000-bit data packet. (c) Overnight (12-hour) shipment of 100 compact disks (650 MB each). 21 Calculate the bandwidth × delay product for the following links. Use one-way delay, measured from first bit sent to first bit received. (a) 10-Mbps Ethernet with a delay of 10 μs. (b) 10-Mbps Ethernet with a single store-and-forward switch like that of Exercise 18(a), packet size 5000 bits, and 10 μs per link propagation delay. (c) 1.5-Mbps T1 link, with a transcontinental one-way delay of 50 ms. (d) 1.5-Mbps T1 link through a satellite in geosynchronous orbit, 35,900 km high. The only delay is speed-of-light propagation delay. 22 Hosts A and B are each connected to a switch S via 10-Mbps links as in Figure 1.25. The propagation delay on each link is 20 μs. S is a store-and- forward device; it begins retransmitting a received packet 35 μs after it has finished receiving it. Calculate the total time required to transmit 10,000 bits from AtoB Exercises 59 ASB Figure 1.25 Diagram for Exercise 22. (a) as a single packet (b) as two 5000-bit packets sent one right after the other 23 Suppose a host has a 1-MB file that is to be sent to another host. The file takes 1 second of CPU time to compress 50%, or 2 seconds to compress 60%. (a) Calculate the bandwidth at which each compression option takes the same total compression + transmission time. (b) Explain why latency does not affect your answer. 24 Suppose that a certain communications protocol involves a per-packet overhead of 100 bytes for headers and framing. We send 1 million bytes of data using this protocol; however, one data byte is corrupted and the entire packet containing it is thus lost. Give the total number of overhead + loss bytes for packet data sizes of 1000, 5000, 10,000, and 20,000 bytes. Which size is optimal? 25 Assume you wish to transfer an n-byte file along a path composed of the source, destination, seven point-to-point links, and five switches. Suppose each link has a propagation delay of 2 ms, bandwidth of 4 Mbps, and that the switches support both circuit and packet switching. Thus you can either break the file up into 1-KB packets, or set up a circuit through the switches and send the file as one contiguous bit stream. Suppose that packets have 24 bytes of packet header information and 1000 bytes of payload, that store-and-forward packet processing at each switch incurs a 1-ms delay after the packet has been completely received, that packets may be sent continuously without waiting for acknowledgments, and that circuit setup requires a 1-KB message to make one round-trip on the path incurring a 1-ms delay at each switch after the message has been completely received. Assume switches introduce no delay to data traversing a circuit. You may also assume that file size is a multiple of 1000 bytes. (a) For what file size n bytes is the total number of bytes sent across the network less for circuits than for packets? (b) For what file size n bytes is the total latency incurred before the entire file arrives at the destination less for circuits than for packets? 60 1 Foundation (c) How sensitive are these results to the number of switches along the path? To the bandwidth of the links? To the ratio of packet size to packet header size? (d) How accurate do you think this model of the relative merits of circuits and packets is? Does it ignore important considerations that discredit one or the other approach? If so, what are they? 26 Consider a closed-loop network (e.g., token ring) with bandwidth 100 Mbps and propagation speed of 2 × 108 m/s. What would the circumference of the loop be to exactly contain one 250-byte packet, assuming nodes do not introduce delay? What would the circumference be if there was a node every 100 m, and each node introduced 10 bits of delay? 27 Compare the channel requirements for voice traffic with the requirements for the real-time transmission of music, in terms of bandwidth, delay, and jitter. What would have to improve? By approximately how much? Could any channel re- quirements be relaxed? 28 For the following, assume that no data compression is done; this would in prac- tice almost never be the case. For (a)–(c), calculate the bandwidth necessary for transmitting in real time: (a) Video at a resolution of 640 × 480, 3 bytes/pixel, 30 frames/second. (b) 160 × 120 video, 1 byte/pixel, 5 frames/second. (c) CD-ROM music, assuming one CD holds 75 minutes’ worth and takes 650 MB. (d) Assume a fax transmits an 8 × 10-inch black-and-white image at a resolution of 72 pixels per inch. How long would this take over a 14.4-Kbps modem? 29 For the following, as in the previous problem, assume that no data compression is done. Calculate the bandwidth necessary for transmitting in real time: (a) HDTV high-definition video at a resolution of 1920 × 1080, 24 bits/pixel, 30 frames/second. (b) POTS (plain old telephone service) voice audio of 8-bit samples at 8 KHz. (c) GSM mobile voice audio of 260-bit samples at 50 Hz. (d) HDCD high-definition audio of 24-bit samples at 88.2 KHz. 30 Discuss the relative performance needs of the following applications, in terms of average bandwidth, peak bandwidth, latency, jitter, and loss tolerance: Exercises 61 (a) File server (b) Print server (c) Digital library (d) Routine monitoring of remote weather instruments (e) Voice (f) Video monitoring of a waiting room (g) Television broadcasting 31 Suppose a shared medium M offers to hosts A1,A2, ...,AN in round-robin fashion an opportunity to transmit one packet; hosts that have nothing to send immediately relinquish M. How does this differ from STDM? How does network utilization of this scheme compare with STDM? 32 Consider a simple protocol for transferring files over a link. After some initial negotiation, A sends data packets of size 1 KB to B; B then replies with an ac- knowledgment. A always waits for each ACK before sending the next data packet; this is known as stop-and-wait. Packets that are overdue are presumed lost and are retransmitted. (a) In the absence of any packet losses or duplications, explain why it is not necessary to include any “sequence number” data in the packet headers. (b) Suppose that the link can lose occasional packets, but that packets that do arrive always arrive in the order sent. Is a 2-bit sequence number (that is, N mod 4) enough for A and B to detect and resend any lost packets? Is a 1-bit sequence number enough? (c) Now suppose that the link can deliver out of order, and that sometimes a packet can be delivered as much as 1 minute after subsequent packets. How does this change the sequence number requirements? 33 Suppose hosts A and B are connected by a link. Host A continuously transmits the current time from a high-precision clock, at a regular rate, fast enough to consume all the available bandwidth. Host B reads these time values and writes them each paired with its own time from a local clock synchronized with A’s. Give qualitative examples of B’s output assuming the link has (a) high bandwidth, high latency, low jitter (b) low bandwidth, high latency, high jitter (c) high bandwidth, low latency, low jitter, occasional lost data 62 1 Foundation For example, a link with zero jitter, a bandwidth high enough to write on every other clock tick, and a latency of 1 tick might yield something like (0000, 0001), (0002, 0003), (0004, 0005). 34 Obtain and build the simplex-talk sample socket program shown in the text. Start one server and one client, in separate windows. While the first client is running, start 10 other clients that connect to the same server; these other clients should most likely be started in the background with their input redirected from a file. What happens to these 10 clients? Do their connect()s fail, or time out, or succeed? Do any other calls block? Now let the first client exit. What happens? Try this with the server value MAX PENDING set to 1 as well. 35 Modify the simplex-talk socket program so that each time the client sends a line to the server, the server sends the line back to the client. The client (and server) will now have to make alternating calls to recv() and send(). 36 Modify the simplex-talk socket program so that it uses UDP as the transport pro- tocol, rather than TCP. You will have to change SOCK STREAM to SOCK DGRAM in both client and server. Then, in the server, remove the calls to listen() and ac- cept(), and replace the two nested loops at the end with a single loop that calls recv() with socket s. Finally, see what happens when two such UDP clients simul- taneously connect to the same UDP server, and compare this to the TCP behavior. 37 Investigate the different options and parameters that you can set for a TCP con- nection. (Do man tcp on Unix.) Experiment with various parameter settings to see how they affect TCP performance. 38 The Unix utility ping can be used to find the RTT to various Internet hosts. Read the man page for ping, and use it to find the RTT to www.cs.princeton.edu in New Jersey and www.cisco.com in California. Measure the RTT values at different times of day, and compare the results. What do you think accounts for the differences? 39 The Unix utility traceroute, or its Windows equivalent tracert, can be used to find the sequence of routers through which a message is routed. Use this to find the path from your site to some others. How well does the number of hops correlate with the RTT times from ping? How well does the number of hops correlate with geographical distance? 40 Use traceroute, above, to map out some of the routers within your organization (or to verify none are used). This Page Intentionally Left Blank Direct Link Networks It is a mistake to look too far ahead. Only one link in the chain of destiny can be handled at a time. —Winston Churchill The simplest network possible is one in which all the hosts are directly connected by some physical medium. This may be a wire or a fiber, and it may cover a small area (e.g., an office building) or a wide area (e.g., transcontinental). Connecting two or more nodes with a suitable medium is only the first step, however. There are five additional problems that must be addressed before the nodes can successfully exchange packets. PROBLEM Physically Connecting Hosts The first is encoding bits onto the wire or fiber so that they can be un- derstood by a receiving host. Second is the matter of delineating the sequence of bits transmitted over the link into complete messages that can be delivered to the end node. This is called the framing problem, and the messages delivered to the end hosts are often called frames. Third, because frames are sometimes corrupted during transmission, it is necessary to detect these errors and take the appropriate action; this is the error detection problem. The fourth issue is making a link appear reliable in spite of the fact that it corrupts frames from time to time. Finally, in those cases where the link is shared by multiple hosts—as opposed to a simple point-to-point link—it is necessary to mediate access to this link. This is the media access control problem. Although these five issues—encoding, framing, error detection, reliable delivery, and access mediation—can be discussed in the abstract, they are very real problems that are addressed in different ways by different networking technologies. This chapter considers these issues in the context of four specific network technologies: point-to- point links, Carrier Sense Multiple Access (CSMA) networks (of which Ethernet is the most famous example), token rings (of which IEEE Standard 802.5 and FDDI are the 2most famous examples), and wireless (for which 802.11 is an emerging standard). The goal of this chapter is simul- taneously to survey the available network technology and to explore these five fundamental issues. Before tackling the specific issues of connecting hosts, this chapter begins by examining the building blocks that will be used: nodes and links. We then explore the first three issues—encoding, framing, and error detection—in the context of a simple point-to-point link. The techniques introduced in these three sections are general and there- fore apply equally well to multiple-access networks. The problem of reliable delivery is considered next. Since link- level reliability is usually not implemented in shared-access networks, this discussion focuses on point-to-point links only. Finally, we address the media access problem in the context of CSMA, token rings, and wireless. Note that these five functions are, in general, imple- mented in a network adaptor—a board that plugs into a host’s I/O bus on one end and into the physical medium on the other end. In other words, bits are exchanged be- tween adaptors, but correct frames are exchanged between nodes. This adaptor is controlled by software running on the node—the device driver—which, in turn, is typically represented as the bottom protocol in a protocol graph. This chapter concludes with a concrete example of a net- work adaptor and sketches the device driver for such an adaptor. 66 2 Direct Link Networks 2.1 Hardware Building Blocks As we saw in Chapter 1, networks are constructed from two classes of hardware building blocks: nodes and links. This statement is just as true for the simplest possible network—one in which a single point-to-point link connects a pair of nodes—as it is for a worldwide internet. This section gives a brief overview of what we mean by nodes and links and, in so doing, defines the underlying technology that we will assume throughout the rest of this book. 2.1.1 Nodes Nodes are often general-purpose computers, like a desktop workstation, a multipro- cessor, or a PC. For our purposes, let’s assume it’s a workstation-class machine. This workstation can serve as a host that users run application programs on, it might be used inside the network as a switch that forwards messages from one link to another, or it might be configured as a router that forwards internet packets from one network to another. In some cases, a network node—most commonly a switch or router inside the network, rather than a host—is implemented by special-purpose hardware. This is usually done for reasons of performance and cost: It is generally possible to build custom hardware that performs a particular function faster and cheaper than a general- purpose processor can perform it. When this happens, we will first describe the basic function being performed by the node as though this function is being implemented in software on a general-purpose workstation, and then explain why and how this functionality might instead be implemented by special hardware. Although we could leave it at that, it is useful to know a little bit about what a workstation looks like on the inside. This information becomes particularly important when we become concerned about how well the network performs. Figure 2.1 gives a simple block diagram of the workstation-class machine we assume throughout this book. There are three key features of this figure that are worth noting. First, the memory on any given machine is finite. It may be 4 MB or it may be 128 MB, but it is not infinite. As pointed out in Section 1.2.2, this is important because memory turns out to be one of the two scarce resources in the network (the other is link bandwidth) that must be carefully managed if we are to provide a fair amount of network capacity to each user. Memory is a scarce resource because on a node that serves as a switch or router, packets must be buffered in memory while waiting their turn to be transmitted over an outgoing link. Second, each node connects to the network via a network adaptor. This adaptor generally sits on the system’s I/O bus and delivers data between the workstation’s memory and the network link. A software module running on the workstation—the device driver—manages this adaptor. It issues commands to the adaptor, telling it, 2.1 Hardware Building Blocks 67 I/O bus (To network) CPU Memory Network adaptorCache Figure 2.1 Example workstation architecture. for example, from what memory location outgoing data should be transmitted and into what memory location incoming data should be stored. Adaptors are discussed in more detail in Section 2.9. Finally, while CPUs are becoming faster at an unbelievable pace, the same is not true of memory. Recent performance trends show processor speeds doubling every 18 months, but memory latency improving at a rate of only 7% each year.The relevance of this difference is that as a network node, a workstation runs at memory speeds, not processor speeds, to a first approximation. This means that the network software needs to be careful about how it uses memory and, in particular, about how many times it accesses memory as it processes each message. We do not have the luxury of being sloppy just because processors are becoming infinitely fast. 2.1.2 Links Network links are implemented on a variety of different physical media, including twisted pair (the wire that your phone connects to), coaxial cable (the wire that your TV connects to), optical fiber (the medium most commonly used for high-bandwidth, long-distance links), and space (the stuff that radio waves, microwaves, and infrared beams propagate through). Whatever the physical medium, it is used to propagate signals. These signals are actually electromagnetic waves traveling at the speed of light. (The speed of light is, however, medium dependent—electromagnetic waves traveling through copper and fiber do so at about two-thirds the speed of light in a vacuum.) 68 2 Direct Link Networks One important property of an electromagnetic wave is the frequency, measured in hertz, with which the wave oscillates. The distance between a pair of adjacent maxima or minima of a wave, typically measured in meters, is called the wave’s wavelength. Since all electromagnetic waves travel at the speed of light, that speed divided by the wave’s frequency is equal to its wavelength. We have already seen the example of a voice-grade telephone line, which carries continuous electromagnetic signals ranging between 300 Hz and 3300 Hz; a 300-Hz wave traveling through copper would have a wavelength of SpeedOfLightInCopper ÷ Frequency = 2/3 × 3 × 108 ÷ 300 = 667 × 103 meters Generally, electromagnetic waves span a much wider range of frequencies, ranging from radio waves, to infrared light, to visible light, to X rays and gamma rays. Figure 2.2 depicts the electromagnetic spectrum and shows which media are commonly used to carry which frequency bands. So far we understand a link to be a physical medium carrying signals in the form of electromagnetic waves. Such links provide the foundation for transmitting all sorts of information, including the kind of data we are interested in transmitting—binary data (1s and 0s). We say that the binary data is encoded in the signal. The problem of encoding binary data onto electromagnetic signals is a complex topic. To help make the topic more manageable, we can think of it as being divided into two layers. The lower layer is concerned with modulation—varying the frequency, amplitude, or phase of the signal to effect the transmission of information. A simple example of modulation Radio Infrared UVMicrowave f(Hz) FM Coax Satellite TV AM Terrestrial microwave Fiber optics X ray 100 104 105 106 107 108 109 1010 1011 1012 1013 1014 1015 1016 102 106 108 1010 1012 1014 1016 1018 1020 1022 1024104 Gamma ray Figure 2.2 Electromagnetic spectrum. 2.1 Hardware Building Blocks 69 is to vary the power (amplitude) of a single wavelength. Intuitively, this is equivalent to turning a light on and off. Because the issue of modulation is secondary to our discussion of links as a building block for computer networks, we simply assume that it is possible to transmit a pair of distinguishable signals—think of them as a “high” signal and a “low” signal—and we consider only the upper layer, which is concerned with the much simpler problem of encoding binary data onto these two signals. Section 2.2 discusses such encodings. Another attribute of a link is how many bit streams can be encoded on it at a given time. If the answer is only one, then the nodes connected to the link must share access to the link. This is the case for the multiple-access links described in Sections 2.6 and 2.7. For point-to-point links, however, it is often the case that two bit streams can be simultaneously transmitted over the link at the same time, one going in each direction. Such a link is said to be full-duplex. A point-to-point link that supports data flowing in only one direction at a time—such a link is called half-duplex—requires that the two nodes connected to the link alternate using it. For the purposes of this book, we assume that all point-to-point links are full-duplex. The only other property of a link that we are interested in at this stage is a very pragmatic one—how do you go about getting one? The answer depends on how far the link needs to reach, how much money you have to spend, and whether or not you know how to operate earth-moving equipment. The following is a survey of different link types you might use to build a computer network. Cables If the nodes you want to connect are in the same room, in the same building, or even on the same site (e.g., a campus), then you can buy a piece of cable and physically string it between the nodes. Exactly what type of cable you choose to install depends on the technology you plan to use to transmit data over the link; we’ll see several examples later in this chapter. For now, a list of the common cable (fiber) types is given in Table 2.1. Of these, Category 5 (Cat-5) twisted pair—it uses a thicker gauge than the twisted pair you find in your home—is quickly becoming the within-building norm. Because of the difficulty and cost in pulling new cable through a building, every effort is made to make new technologies use existing cable; Gigabit Ethernet, for example, has been designed to run over Cat-5 wiring. Fiber is typically used to connect buildings at a site. Leased Lines If the two nodes you want to connect are on opposite sides of the country, or even across town, then it is not practical to install the link yourself. Your only option is to lease a dedicated link from the telephone company, in which case all you’ll need to be able to do is conduct an intelligent conversation with the phone company customer 70 2 Direct Link Networks Cable Typical Bandwidths Distances Category 5 twisted pair 10–100 Mbps 100 m Thin-net coax 10–100 Mbps 200 m Thick-net coax 10–100 Mbps 500 m Multimode fiber 100 Mbps 2km Single-mode fiber 100–2400 Mbps 40 km Table 2.1 Common types of cables and fibers available for local links. Service Bandwidth DS1 1.544 Mbps DS3 44.736 Mbps STS-1 51.840 Mbps STS-3 155.250 Mbps STS-12 622.080 Mbps STS-48 2.488320 Gbps STS-192 9.953280 Gbps Table 2.2 Common bandwidths available from the carriers. service representative. Table 2.2 gives the common services that can be leased from the phone company. Again, more details are given throughout this chapter. While these bandwidths appear somewhat arbitrary, there is actually some method to the madness. DS1 and DS3 (they are also sometimes called T1 and T3, respectively) are relatively old technologies that were orginally defined for copper- based transmission media. DS1 is equal to the aggregation of 24 digital voice circuits of 64 Kbps each, and DS3 is equal to 28 DS1 links. All the STS-N links are for optical fiber (STS stands for Synchronous Transport Signal). STS-1 is the base link speed, and each STS-N has N times the bandwidth of STS-1. An STS-N link is also sometimes called an OC-N link (OC stands for optical carrier). The difference between STS and OC is subtle: The former refers to the electrical transmission on the devices connected to the link, and the latter refers to the actual optical signal that is propagated over the fiber. Keep in mind that the phone company does not implement the “link” we just ordered as a single, unbroken piece of cable or fiber. Instead, it implements the link 2.1 Hardware Building Blocks 71 on its own network. Although the telephone network has historically looked much different from the kind of network described in this book—it was built primarily to provide a voice service and used circuit-switching technology—the current trend is toward the style of networking described in this book, including the asynchronous transfer mode (ATM) network described in Chapter 3. This is not surprising—the potential market for carrying data, voice, and video is huge. In any case, whether the link is physical or a logical connection through the telephone network, the problem of building a computer network on top of a collection of such links remains the same. So, we will proceed as though each link is implemented by a single cable/fiber, and only when we are done will we worry about whether we have just built a computer network on top of the underlying telephone network, or the computer network we have just built could itself serve as the backbone for the telephone network. Last-Mile Links If you can’t afford a dedicated leased line—they range in price from roughly a thousand dollars a month for a cross-country DS1 link to “if you have to ask, you can’t afford it”—then there are less expensive options available. We call these “last-mile” links because they often span the last mile from the home to a network service provider. These services, which are summarized in Table 2.3, typically connect a home to an existing network. This means they are probably not suitable for use in building a com- plete network from scratch, but if you’ve already succeeded in building a network—and “you” happen to be either the telephone company or the cable company—then you can use these links to reach millions of customers. The first option is a conventional modem over POTS (plain old telephone ser- vice). Today it is possible to buy a modem that transmits data at 56 Kbps over a standard voice-grade line for less than a hundred dollars. The technology is already at its bandwidth limit, however, which has led to the development of the second option: ISDN (Integrated Services Digital Network). An ISDN connection includes two Service Bandwidth POTS 28.8–56 Kbps ISDN 64–128 Kbps xDSL 16 Kbps–55.2 Mbps CATV 20–40 Mbps Table 2.3 Common services available to connect your home. 72 2 Direct Link Networks 1.554–8.448 Mbps 16–640 Kbps Local loop Central office Subscriber premises Figure 2.3 ADSL connects the subscriber to the central office via the local loop. 64-Kbps channels, one that can be used to transmit data and another that can be used for digitized voice. (A device that encodes analog voice into a digital ISDN link is called a CODEC, for coder/decoder.) When the voice channel is not in use, it can be combined with the data channel to support up to 128 Kbps of data bandwidth. For many years ISDN was viewed as the future for modest bandwidth into the home. ISDN has now been largely overtaken, however, by two newer tech- nologies: xDSL (digital subscriber line) and cable modems. The former is actually a collection of technologies that are able to transmit data at high speeds over the standard twisted pair lines that currently come into most homes in the United States (and many other places). The one in most widespread use today is ADSL (asymmetric digital subscriber line). As its name implies, ADSL provides a different bandwidth from the subscriber to the telephone company’s central office (upstream) than it does from the central office to the subscriber (down- stream). The exact bandwidth depends on the length of the line running from the subscriber to the central office. This line is called the local loop, as illustrated in Figure 2.3, and runs over existing copper. Downstream bandwidths range from 1.544 Mbps (18,000 feet) to 8.448 Mbps (9000 feet), while upstream bandwidths range from 16 Kbps to 640 Kbps. Shannon’s Theorem Meets Your Modem There has been an enormous body of work done in the related areas of signal processing and informa- tion theory, studying everything from how signals degrade over dis- tance to how much data a given sig- nal can effectively carry. The most notable piece of work in this area is a formula known as Shannon’s theorem. Simply stated, Shannon’s theorem gives an upper bound to the capacity of a link, in terms of bits per second (bps), as a function of the signal-to-noise ratio of the link, measured in decibels (dB). Shannon’s theorem can be used to determine the data rate at which a modem can be expected to trans- mit binary data over a voice-grade phone line without suffering from 2.1 Hardware Building Blocks 73 STS-N over fiber VDSL at 12.96–55.2 Mbps over 1000–4500 feet of copper Central office Subscriber premises Neighborhood optical network unit Figure 2.4 VDSL connects the subscriber to the optical network that reaches the neighborhood. An alternative technology that has yet to be widely deployed—very high data rate digital subscriber line (VDSL)—is symmetric, with data rates ranging from 12.96 Mbps to 55.2 Mbps. VDSL runs over much shorter distances—1000 to 4500 feet—which means that it will not typically reach from the home to the central office. Instead, the telephone company would have to put VDSL transmission hardware in neighbor- hoods, with some other technology (e.g., STS-N running over fiber) connecting the too high an error rate. For example, we assume that a voice-grade phone connection supports a fre- quency range of 300 Hz to 3300 Hz. Shannon’s theorem is typically given by the following formula: C = B log2(1 + S/N) where C is the achievable channel capacity measured in hertz, B is the bandwidth of the line (3300 Hz −300 Hz = 3000 Hz), S is the av- erage signal power, and N is the average noise power. The signal- to-noise ratio (S/N) is usually expressed in decibels, related as fol- lows: dB = 10 × log10(S/N) neighborhood to the central office, as illus- trated in Figure 2.4. This is sometimes called “fiber to the neighborhood” (contrasting with more ambitious schemes such as “fiber to the home” and “fiber to the curb”). Cable modems are an alternative to the various types of DSL. As the name suggests, this technology uses the cable TV (CATV) infrastructure, which currently reaches 95% of the households in the United States. (Only 65% of U.S. homes actually subscribe.) In this approach, some subset of the available CATV channels are made available for transmitting dig- ital data, where a single CATV channel has a bandwidth of 6 MHz. CATV, like ADSL, is used in an asymmetric way, with downstream rates much greater than up- stream rates. The technology is currently able to achieve 40 Mbps downstream on a single CATV channel, with 100 Mbps as the theoretical capacity. The upstream rate is roughly half the downstream rate (i.e., 20 Mbps) due to a 1000-fold de- crease in the signal-to-noise ratio. It is also the case that fewer CATV channels are 74 2 Direct Link Networks dedicated to upstream traffic than to downstream traffic. Unlike DSL, the bandwidth is shared among all subscribers in a neighborhood (a fact that led to some amusing advertising from DSL providers). This means that some method for arbitrating access to the shared medium—similar to the 802 standards described later in this chapter— needs to be used. Finally, like DSL, it is unlikely that cable modems will be used to connect arbitrary node A at one site to arbitrary node B at some other site. Instead, cable modems are seen as a means to connect node A in your home to the cable company, with the cable company then defining what the rest of the network looks like. Wireless Links The field of wireless communication is exploding, both economically and technologi- cally. The Advanced Mobile Phone System (AMPS) has been the standard for cellular phones in the United States for several years. AMPS, which is based on analog technol- ogy, is rapidly giving way to digital cellular–PCS (Personal Communication Services) in the United States and Canada, and GSM (Global System for Mobile Communi- cation) in the rest of the world. All three systems currently use a system of towers to transmit signals, although some signifi- cant efforts have been made to supplement this infrastructure by ringing the globe with a grid of medium- and low-orbit satellites. These projects—which include ICO, Glob- alstar, Iridium, and Teledesic—have had mixed success. Those that are still viable are mostly focusing on delivery of telephone ser- vice to those increasingly rare parts of the globe where cellular service is not available. Thinking a bit less globally, frequency bands from the radio and infrared portions of the electromagnetic spectrum can be used to provide wireless links over short dis- tances, such as inside office buildings, cof- fee shops, building complexes, and cam- puses. In the case of infrared, signals with wavelengths in the 850–950-nanometer range can be used to transmit data at 1-Mbps rates over distances of about 10 m. Assuming a typical decibel ratio of 30 dB, this means that S/N =1000. Thus, we have C = 3000 × log2(1001) which equals approximately 30 Kbps, roughly the limit of a 28.8- Kbps modem. Given this fundamental limit, why is it possible to buy 56-Kbps modems at any electronics store? One reason is that such rates de- pend on improved line quality, that is, a higher signal-to-noise ratio than 30 dB. Another reason is that changes within the phone sys- tem have largely eliminated analog lines that are bandwidth-limited to 3300 Hz. 2.2 Encoding (NRZ, NRZI, Manchester, 4B/5B) 75 This technology does not require line of sight, but is limited to in-building environ- ments. In the case of radio, several different bands are currently being made available for data communication. For example, bands at 5.2 GHz and 17 GHz are allocated to HIPERLAN (High Performance European Radio LAN) in Europe. Similarly, band- width at 2.4 GHz has been set aside in many countries for use with the IEEE 802.11 standard for wireless LANs. (Additional bandwidth is available at 5 GHz, but unfor- tunately it is subject to interference from microwave ovens.) IEEE 802.11, which is an evolving standard that supports data rates of up to 54 Mbps, will be discussed more fully in Section 2.8. Another interesting development in the wireless arena is the Bluetooth radio interface that operates in the 2.45-GHz frequency band. Bluetooth is designed for short distances (on the order of 10 m) with a bandwidth of 1 Mbps. Its developers envision it being used in all devices (e.g., printers, workstations, laptops, projectors, PDAs, mobile phones), thereby eliminating the need for wires and cables in the office (or between the various devices on your body, perhaps). Networks of such devices are starting to be called piconets. 2.2 Encoding (NRZ, NRZI, Manchester, 4B/5B) The first step in turning nodes and links into usable building blocks is to understand how to connect them in such a way that bits can be transmitted from one node to the other. As mentioned in the preceding section, signals propagate over physical links. The task, therefore, is to encode the binary data that the source node wants to send into the signals that the links are able to carry, and then to decode the signal back into the corresponding binary data at the receiving node. We ignore the details of modulation and assume we are working with two discrete signals: high and low. In practice, these signals might correspond to two different voltages on a copper-based link, or two different power levels on an optical link. As we have said, most of the functions discussed in this chapter are performed by a network adaptor—a piece of hardware that connects a node to a link. The network adaptor contains a signalling component that actually encodes bits into signals at the sending node and decodes signals into bits at the receiving node. Thus, as illustrated in Figure 2.5, signals travel over a link between two signalling components, and bits flow between network adaptors. Let’s return to the problem of encoding bits onto signals. The obvious thing to do is to map the data value 1 onto the high signal and the data value 0 onto the low signal. This is exactly the mapping used by an encoding scheme called, cryptically enough, non-return to zero (NRZ). For example, Figure 2.6 schematically depicts the NRZ-encoded signal (bottom) that corresponds to the transmission of a particular sequence of bits (top). 76 2 Direct Link Networks Signalling component Signal Bits Node NodeAdaptor Adaptor Figure 2.5 Signals travel between signalling components; bits flow between adaptors. Bits NRZ 0010111101000010 Figure 2.6 NRZ encoding of a bit stream. The problem with NRZ is that a sequence of several consecutive 1s means that the signal stays high on the link for an extended period of time, and similarly, sev- eral consecutive 0s means that the signal stays low for a long time. There are two fundamental problems caused by long strings of 1s or 0s. The first is that it leads to a situation known as baseline wander. Specifically, the receiver keeps an average of the signal it has seen so far,and then uses this average to distinguish between low and high signals. Whenever the signal is significantly lower than this average, the receiver con- cludes that it has just seen a 0, and likewise, a signal that is significantly higher than the average is interpreted to be a 1. The prob- lem, of course, is that too many consecutive 1s or 0s cause this average to change, mak- ing it more difficult to detect a significant change in the signal. The second problem is that frequent transitions from high to low and vice versa are necessary to enable clock recovery. Intuitively, the clock recovery problem is that both the encoding and the decoding Bit Rates and Baud Rates Many people use the terms bit rate and baud rate interchangeably, even though as we see with the Manchester encoding, they are not the same thing. While the Man- chester encoding is an example of a case in which a link’s baud rate is greater than its bit rate, it is also possible to have a bit rate that is greater than the baud rate. This would imply that more than one bit is encoded on each pulse sent over the link. To see how this might happen, suppose you could transmit four 2.2 Encoding (NRZ, NRZI, Manchester, 4B/5B) 77 Bits NRZ Clock Manchester NRZI 0010111101000010 Figure 2.7 Different encoding strategies. processes are driven by a clock—every clock cycle the sender transmits a bit and the receiver recovers a bit. The sender’s and the receiver’s clocks have to be precisely synchronized in order for the receiver to recover the same bits the sender trans- mits. If the receiver’s clock is even slightly faster or slower than the sender’s clock, then it does not correctly decode the signal. You could imagine sending the clock to the receiver over a separate wire, but this is typically avoided because it makes the cost of cabling twice as high. So instead, the receiver derives the clock from the distinguished signals over a link rather than just two. On an analog link, for example, these four signals might correspond to four different frequencies. Given four different signals, it is possible to encode two bits of information on each signal. That is, the first signal means 00, the second signal means 01, and so on. Now, a sender (receiver) that is able to transmit (detect) 1000 pulses per second would be able to send (receive) 2000 bits of informa- tion per second. That is, it would be a 1000-baud/2000-bps link. received signal—the clock recovery process. Whenever the signal changes, such as on a transition from 1 to 0 or from 0 to 1, then the receiver knows it is at a clock cycle boundary, and it can resynchronize itself. However, a long period of time with- out such a transition leads to clock drift. Thus, clock recovery depends on having lots of transitions in the signal, no matter what data is being sent. One approach that addresses this problem, called non-return to zero inverted (NRZI), has the sender make a transition from the current signal to encodea1and stay at the current signal to encode a 0. This solves the problem of consecutive 1s, but obviously does nothing for consecu- tive 0s. NRZI is illustrated in Figure 2.7. An alternative, called Manchester encod- ing, does a more explicit job of merging 78 2 Direct Link Networks the clock with the signal by transmitting the exclusive-OR of the NRZ-encoded data and the clock. (Think of the local clock as an internal signal that alternates from low to high; a low/high pair is considered one clock cycle.) The Manchester encoding is also illustrated in Figure 2.7. Observe that the Manchester encoding results in 0 being encoded as a low-to-high transition and 1 being encoded as a high-to-low transition. Because both 0s and 1s result in a transition to the signal, the clock can be effectively recovered at the receiver. (There is also a variant of the Manchester encoding, called differential Manchester, in whicha1isencoded with the first half of the signal equal to the last half of the previous bit’s signal anda0isencoded with the first half of the signal opposite to the last half of the previous bit’s signal.) The problem with the Manchester encoding scheme is that it doubles the rate at which signal transitions are made on the link, which means that the receiver has half the time to detect each pulse of the signal. The rate at which the signal changes is called the link’s baud rate. In the case of the Manchester encoding, the bit rate is half the baud rate, so the encoding is considered only 50% efficient. Keep in mind that if the receiver had been able to keep up with the faster baud rate required by the Manchester encoding in Figure 2.7, then both NRZ and NRZI could have been able to transmit twice as many bits in the same time period. A final encoding that we consider, called 4B/5B, attempts to address the inef- ficiency of the Manchester encoding without suffering from the problem of having extended durations of high or low signals. The idea of 4B/5B is to insert extra bits into the bit stream so as to break up long sequences of 0s or 1s. Specifically, every 4 bits of actual data are encoded in a 5-bit code that is then transmitted to the receiver; hence the name 4B/5B. The 5-bit codes are selected in such a way that each one has no more than one leading 0 and no more than two trailing 0s. Thus, when sent back-to-back, no pair of 5-bit codes results in more than three consecutive 0s being transmitted. The resulting 5-bit codes are then transmitted using the NRZI encoding, which explains why the code is only concerned about consecutive 0s—NRZI already solves the problem of consecutive 1s. Note that the 4B/5B encoding results in 80% efficiency. Table 2.4 gives the 5-bit codes that correspond to each of the 16 possible 4-bit data symbols. Notice that since 5 bits are enough to encode 32 different codes, and we are using only 16 of these for data, there are 16 codes left over that we can use for other purposes. Of these, code 11111 is used when the line is idle, code 00000 corresponds to when the line is dead, and 00100 is interpreted to mean halt. Of the remaining 13 codes, 7 of them are not valid because they violate the “one leading 0, two trailing 0s” rule, and the other 6 represent various control symbols. As we will see later in this chapter, some framing protocols (e.g., FDDI) make use of these control symbols. 2.3 Framing 79 4-Bit Data Symbol 5-Bit Code 0000 11110 0001 01001 0010 10100 0011 10101 0100 01010 0101 01011 0110 01110 0111 01111 1000 10010 1001 10011 1010 10110 1011 10111 1100 11010 1101 11011 1110 11100 1111 11101 Table 2.4 4B/5B encoding. 2.3 Framing Now that we have seen how to transmit a sequence of bits over a point-to-point link— from adaptor to adaptor—let’s consider the scenario illustrated in Figure 2.8. Recall from Chapter 1 that we are focusing on packet-switched networks, which means that blocks of data (called frames at this level), not bit streams, are exchanged between nodes. It is the network adaptor that enables the nodes to exchange frames. When node A wishes to transmit a frame to node B, it tells its adaptor to transmit a frame from the node’s memory. This results in a sequence of bits being sent over the link. The adaptor on node B then collects together the sequence of bits arriving on the link and deposits the corresponding frame in B’s memory. Recognizing exactly what set of bits constitutes a frame—that is, determining where the frame begins and ends—is the central challenge faced by the adaptor. 80 2 Direct Link Networks Frames Bits Node A Node BAdaptor Adaptor Figure 2.8 Bits flow between adaptors, frames between hosts. SYN Header Body 88 8 8168 SYN SOH STX ETX CRC Figure 2.9 BISYNC frame format. There are several ways to address the framing problem. This section uses several different protocols to illustrate the various points in the design space. Note that while we discuss framing in the context of point-to-point links, the problem is a fundamental one that must also be addressed in multiple-access networks like Ethernet and token rings. 2.3.1 Byte-Oriented Protocols (BISYNC, PPP, DDCMP) One of the oldest approaches to framing—it has its roots in connecting terminals to mainframes—is to view each frame as a collection of bytes (characters) rather than a collection of bits. Such a byte-oriented approach is exemplified by the BISYNC (Binary Synchronous Communication) protocol developed by IBM in the late 1960s, and the DDCMP (Digital Data Communication Message Protocol) used in Digital Equipment Corporation’s DECNET. Sometimes these protocols assume a particular character set—for example, BISYNC can support ASCII, EBCDIC, and IBM’s 6-bit Transcode—but this is not necessarily the case. Although similar in many respects, these two protocols are examples of two different framing techniques, the sentinel approach and the byte-counting approach. Sentinel Approach The BISYNC protocol illustrates the sentinel approach to framing; its frame format is depicted in Figure 2.9. This figure is the first of many that you will see in this book that are used to illustrate frame or packet formats, so a few words of explanation 2.3 Framing 81 are in order. We show a packet as a sequence of labeled fields. Above each field is a number indicating the length of that field in bits. Note that the packets are transmitted beginning with the leftmost field. The beginning of a frame is denoted by sending a special SYN (synchronization) character. The data portion of the frame is then contained between special sentinel characters: STX (start of text) and ETX (end of text). The SOH (start of header) field serves much the same purpose as the STX field. The problem with the sentinel approach, of course, is that the ETX character might appear in the data portion of the frame. BISYNC overcomes this problem by “escaping” the ETX character by preceding it with a DLE (data-link-escape) character whenever it appears in the body of a frame; the DLE character is also escaped (by preceding it with an extra DLE) in the frame body. (C programmers may notice that this is analogous to the way a quotation mark is escaped by the backslash when it occurs inside a string.) This approach is often called character stuffing because extra characters are inserted in the data portion of the frame. The frame format also includes a field labeled CRC (cyclic redundancy check) that is used to detect transmission errors; various algorithms for error detection are presented in Section 2.4. Finally, the frame contains additional header fields that are used for, among other things, the link-level reliable delivery algorithm. Examples of these algorithms are given in Section 2.5. The more recent Point-to-Point Protocol (PPP), which is commonly run over dial- up modem links, is similar to BISYNC in that it uses character stuffing. The format for a PPP frame is given in Figure 2.10. The special start-of-text character, denoted as the Flag field in Figure 2.10, is 01111110. The Address and Control fields usually contain default values, and so are uninteresting. The Protocol field is used for demultiplexing: It identifies the high-level protocol such as IP or IPX (an IP-like protocol developed by Novell). The frame payload size can be negotiated, but it is 1500 bytes by default. The Checksum field is either 2 (by default) or 4 bytes long. The PPP frame format is unusual in that several of the field sizes are negotiated rather than fixed. This negotiation is conducted by a protocol called LCP (Link Control Protocol). PPP and LCP work in tandem: LCP sends control messages encapsulated in PPP frames—such messages are denoted by an LCP identifier in the PPP Protocol ProtocolControlAddressFlag Payload 88816168 FlagChecksum Figure 2.10 PPP frame format. 82 2 Direct Link Networks SYN Header Body 88 4214 168 SYN Class CRCCount Figure 2.11 DDCMP frame format. field—and then turns around and changes PPP’s frame format based on the information contained in those control messages. LCP is also involved in establishing a link between two peers when both sides detect the carrier signal. Byte-Counting Approach As every Computer Sciences 101 student knows, the alternative to detecting the end of a file with a sentinel value is to include the number of items in the file at the beginning of the file. The same is true in framing—the number of bytes contained in a frame can be included as a field in the frame header. DECNET’s DDCMP protocol uses this approach, as illustrated in Figure 2.11. In this example, the COUNT field specifies how many bytes are contained in the frame’s body. One danger with this approach is that a transmission error could corrupt the COUNT field, in which case the end of the frame would not be correctly detected. (A similar problem exists with the sentinel- based approach if the ETX field becomes corrupted.) Should this happen, the receiver will accumulate as many bytes as the bad COUNT field indicates and then use the error detection field to determine that the frame is bad. This is sometimes called a framing error. The receiver will then wait until it sees the next SYN character to start collecting the bytes that make up the next frame. It is What’s in a Layer? One of the important contribu- tions of the OSI reference model presented in Chapter 1 was to pro- vide some vocabulary for talking about protocols and, in particular, protocol layers. This vocabulary has provided fuel for plenty of arguments along the lines of “Your protocol does function X at layer Y, and the OSI reference model says it should be done at layer Z—that’s a layer violation.” In fact, figuring out the right layer at which to perform a given function can be very difficult, and the reasoning is usually a lot more subtle than “What does the OSI model say?” It is partly for this reason that this book avoids a rigidly layerist approach. Instead, it shows you a lot of functions that need to be performed by protocols and looks at some ways that they have been successfully implemented. 2.3 Framing 83 Header Body 816 16 8 CRCBeginning sequence Ending sequence Figure 2.12 HDLC frame format. therefore possible that a framing error will cause back-to-back frames to be incorrectly received. In spite of our nonlayerist approach, sometimes we need con- venient ways to talk about classes of protocols, and the name of the layer at which they operate is often the best choice. Thus, for exam- ple, this chapter focuses primar- ily on link-layer protocols. (Bit encoding, described in Section 2.2, is the exception, being considered a physical-layer function.) Link- layer protocols can be identified by the fact that they run over sin- gle links—the type of network dis- cussed in this chapter. Network- layer protocols, by contrast, run over switched networks that con- tain lots of links interconnected by switches or routers. Topics related to network-layer protocols are dis- cussed in Chapters 3 and 4. Note that protocol layers are supposed to be helpful—they pro- vide helpful ways to talk about 2.3.2 Bit-Oriented Protocols (HDLC) Unlike these byte-oriented protocols, a bit- oriented protocol is not concerned with byte boundaries—it simply views the frame as a collection of bits. These bits might come from some character set, such as ASCII, they might be pixel values in an image, or they could be instructions and operands from an executable file. The Synchronous Data Link Control (SDLC) protocol developed by IBM is an example of a bit-oriented protocol; SDLC was later standardized by the ISO as the High-Level Data Link Control (HDLC) protocol. In the following discussion, we use HDLC as an example; its frame format is given in Figure 2.12. HDLC denotes both the beginning and the end of a frame with the distin- guished bit sequence 01111110. This sequ- ence is also transmitted during any times that the link is idle so that the sender and receiver can keep their clocks synchro- nized. In this way, both protocols essentially use the sentinel approach. Because this se- quence might appear anywhere in the body of the frame—in fact, the bits 01111110 84 2 Direct Link Networks might cross byte boundaries—bit-oriented protocols use the analog of the DLE char- acter, a technique known as bit stuffing. Bit stuffing in the HDLC protocol works as follows. On the sending side, any time five consecutive 1s have been transmitted from the body of the message (i.e., excluding when the sender is trying to transmit the distinguished 01111110 sequence), the sender inserts a 0 before transmitting the next bit. On the receiving side, should five consecutive 1s arrive, the receiver makes its decision based on the next bit it sees (i.e., the bit following the five 1s). If the next bit is a 0, it must have been stuffed, and so the receiver removes it. If the next bit is a 1, then one of two things is true: Either this is the end-of-frame marker or an error has been introduced into the bit stream. By looking at the next bit, the receiver can distinguish between these two cases: If it sees a 0 (i.e., the last eight bits it has looked at are 01111110), then it is the end-of- frame marker; if it sees a 1 (i.e., the last eight bits it has looked at are 01111111), then there must have been an error and the whole frame is discarded. In the latter case, the receiver has to wait for the next 01111110 before it can start receiving again, and as a consequence, there is the potential that the receiver will fail to receive two consecutive frames. Obviously, there are still ways that framing errors can go undetected, such as when an entire spurious end-of-frame pattern is generated by errors, but these failures are relatively unlikely. Robust ways of detecting errors are discussed in Section 2.4. An interesting characteristic of bit stuffing, as well as character stuffing, is that the size of a frame is dependent on the data that is being sent in the payload of the frame. It is in fact not possible to make all frames exactly the same size, given that the data that might be carried in any frame is arbi- trary. (To convince yourself of this, consider what happens if the last byte of a frame’s body is the ETX character.) A form of fram- ing that ensures that all frames are the same size is described in the next subsection. 2.3.3 Clock-Based Framing (SONET) A third approach to framing is exempli- fied by the Synchronous Optical Network (SONET) standard. For lack of a widely accepted generic term, we refer to this approach simply as clock-based framing. classes of protocols, and they help us divide the problem of build- ing networks into manageable sub- tasks. However, they are not meant to be overly restrictive—the mere fact that something is a layer vio- lation does not end the argument about whether it is a worthwhile thing to do. In other words, lay- ering makes a good slave, but a poor master.A particularly interest- ing argument about the best layer to place a certain function comes up when we look at congestion control in Chapter 6. 2.3 Framing 85 SONET was first proposed by Bell Communications Research (Bellcore), and then developed under the American National Standards Institute (ANSI) for digital trans- mission over optical fiber; it has since been adopted by the ITU-T. Who standardized what and when is not the interesting issue, though. The thing to remember about SONET is that it is the dominant standard for long-distance transmission of data over optical networks. An important point to make about SONET before we go any further is that the full specification is substantially larger than this book. Thus, the following discussion will necessarily cover only the high points of the standard. Also, SONET addresses both the framing problem and the encoding problem. It also addresses a problem that is very important for phone companies—the multiplexing of several low-speed links onto one high-speed link. We begin with framing and discuss the other issues following. As with the previously discussed framing schemes, a SONET frame has some special information that tells the receiver where the frame starts and ends. However, that is about as far as the similarities go. Notably, no bit stuffing is used, so that a frame’s length does not depend on the data being sent. So the question to ask is, How does the receiver know where each frame starts and ends? We consider this question for the lowest-speed SONET link, which is known as STS-1 and runs at 51.84 Mbps. An STS-1 frame is shown in Figure 2.13. It is arranged as nine rows of 90 bytes each, and the first 3 bytes of each row are overhead, with the rest being available for data that is being transmitted over the link. The first 2 bytes of the frame contain a special bit pattern, and it is these bytes that enable the receiver to determine where the frame starts. However, since bit stuffing is not used, there is no reason why this pattern will not occasionally turn up in the payload portion of the frame. To guard against this, the receiver looks for the special bit pattern consistently, hoping to see it appearing Overhead Payload 90 columns 9 rows Figure 2.13 A SONET STS-1 frame. 86 2 Direct Link Networks once every 810 bytes, since each frame is 9 × 90 = 810 bytes long. When the special pattern turns up in the right place enough times, the receiver concludes that it is in sync and can then interpret the frame correctly. One of the things we are not describing due to the complexity of SONET is the detailed use of all the other overhead bytes. Part of this complexity can be attributed to the fact that SONET runs across the carrier’s optical network, not just over a single link. (Recall that we are glossing over the fact that the carriers implement a network, and we are instead focusing on the fact that we can lease a SONET link from them and then use this link to build our own packet-switched network.) Additional complexity comes from the fact that SONET provides a considerably richer set of services than just data transfer. For example, 64 Kbps of a SONET link’s capacity is set aside for a voice channel that is used for maintenance. The overhead bytes of a SONET frame are encoded using NRZ, the simple encoding described in the previous section where 1s are high and 0s are low. However, to ensure that there are plenty of transitions to allow the receiver to recover the sender’s clock, the payload bytes are scrambled. This is done by calculating the exclusive-OR (XOR) of the data to be transmitted and by the use of a well-known bit pattern. The bit pattern, which is 127 bits long, has plenty of transitions from 1 to 0, so that XORing it with the transmitted data is likely to yield a signal with enough transitions to enable clock recovery. SONET supports the multiplexing of multiple low-speed links in the following way. A given SONET link runs at one of a finite set of possible rates, ranging from 51.84 Mbps (STS-1) to 2488.32 Mbps (STS-48) and beyond. (See Table 2.2 in Section 2.1 for the full set of SONET data rates.) Note that all of these rates are integer multiples of STS-1. The significance for framing is that a single SONET frame can contain subframes for multiple lower-rate channels. A second related feature is that each frame is 125 μs long. This means that at STS-1 rates, a SONET frame is 810 bytes long, while at STS-3 rates, each SONET frame is 2430 bytes long. Notice the synergy between these two features: 3 × 810 = 2430, meaning that three STS-1 frames fit exactly in a single STS-3 frame. Intuitively, the STS-N frame can be thought of as consisting of N STS-1 frames, where the bytes from these frames are interleaved; that is, a byte from the first frame is transmitted, then a byte from the second frame is transmitted, and so on. The reason for interleaving the bytes from each STS-N frame is to ensure that the bytes in each STS-1 frame are evenly paced; that is, bytes show up at the receiver at a smooth 51 Mbps, rather than all bunched up during one particular 1/Nth of the 125-μs interval. Although it is accurate to view an STS-Nsignal as being used to multiplex NSTS- 1 frames, the payload from these STS-1 frames can be linked together to form a larger STS-N payload; such a link is denoted STS-Nc (for concatenated). One of the fields in 2.3 Framing 87 STS-1 Hdr STS-1 Hdr STS-1 Hdr STS-3cHdr Figure 2.14 Three STS-1 frames multiplexed onto one STS-3c frame. Frame 0 Frame 1 87 columns 9 rows Figure 2.15 SONET frames out of phase. the overhead is used for this purpose. Figure 2.14 schematically depicts concatenation in the case of three STS-1 frames being concatenated into a single STS-3c frame. The significance of a SONET link being designated as STS-3c rather than STS-3 is that, in the former case, the user of the link can view it as a single 155.25-Mbps pipe, whereas an STS-3 should really be viewed as three 51.84-Mbps links that happen to share a fiber. Finally, the preceding description of SONET is overly simplistic in that it assumes that the payload for each frame is completely contained within the frame. (Why wouldn’t it be?) In fact, we should view the STS-1 frame just described as simply a placeholder for the frame, where the actual payload may float across frame bound- aries. This situation is illustrated in Figure 2.15. Here we see both the STS-1 payload floating across two STS-1 frames, and the payload shifted some number of bytes to the right and, therefore, wrapped around. One of the fields in the frame overhead points to the beginning of the payload. The value of this capability is that it simplifies the task of synchronizing the clocks used throughout the carriers’ networks, which is something that carriers spend a lot of their time worrying about. 88 2 Direct Link Networks 2.4 Error Detection As discussed in Chapter 1, bit errors are sometimes introduced into frames. This happens, for example, because of electrical interference or thermal noise. Although errors are rare, especially on optical links, some mechanism is needed to detect these errors so that corrective action can be taken. Otherwise, the end user is left wondering why the C program that successfully compiled just a moment ago now suddenly has a syntax error in it, when all that happened in the interim is that it was copied across a network file system. There is a long history of techniques for dealing with bit errors in computer systems, dating back to Hamming and Reed/Solomon codes that were developed for use when storing data on magnetic disks and in early core memories. This sec- tion describes some of the error detection techniques most commonly used in net- working. Detecting errors is only one part of the problem. The other part is correcting errors once detected. There are two basic approaches that can be taken when the recipient of a message detects an error. One is to notify the sender that the message was corrupted so that the sender can retransmit a copy of the message. If bit errors are rare, then in all probability the retransmitted copy will be error-free. Alternatively, there are some types of error detection algorithms that allow the recipient to recon- struct the correct message even after it has been corrupted; such algorithms rely on error-correcting codes, discussed below. One of the most common techniques for detecting transmission errors is a tech- nique known as the cyclic redundancy check (CRC). It is used in nearly all the link-level protocols discussed in the previous section—for example, HDLC, DDCMP—as well as in the CSMA and token ring protocols described later in this chapter. Section 2.4.3 outlines the basic CRC algorithm. Before discussing that approach, we consider two simpler schemes that are also widely used: two-dimensional parity and checksums. The former is used by the BISYNC protocol when it is transmitting ASCII characters (CRC is used as the error code when BISYNC is used to transmit EBCDIC), and the latter is used by several Internet protocols. The basic idea behind any error detection scheme is to add redundant information to a frame that can be used to determine if errors have been introduced. In the extreme, we could imagine transmitting two complete copies of the data. If the two copies are identical at the receiver, then it is probably the case that both are correct. If they differ, then an error was introduced into one (or both) of them, and they must be discarded. This is a rather poor error detection scheme for two reasons. First, it sends n redundant bits for an n-bit message. Second, many errors will go undetected—any error that happens to corrupt the same bit positions in the first and second copies of the message. 2.4 Error Detection 89 Fortunately, we can do a lot better than this simple scheme. In general, we can provide quite strong error detection capability while sending only k redundant bits for an n-bit message, where k ≪ n. On an Ethernet, for example, a frame carrying up to 12,000 bits (1500 bytes) of data requires only a 32-bit CRC code, or as it is commonly expressed, uses CRC-32. Such a code will catch the overwhelming majority of errors, as we will see below. We say that the extra bits we send are redundant because they add no new information to the message. Instead, they are derived directly from the original mes- sage using some well-defined algorithm. Both the sender and the receiver know exactly what that algorithm is. The sender applies the algorithm to the message to generate the redundant bits. It then transmits both the message and those few extra bits. When the receiver applies the same algorithm to the received message, it should (in the absence of errors) come up with the same result as the sender. It compares the result with the one sent to it by the sender. If they match, it can conclude (with high likelihood) that no errors were introduced in the message during transmission. If they do not match, it can be sure that either the message or the redundant bits were corrupted, and it must take appropriate action, that is, discarding the message, or correcting it if that is possible. One note on the terminology for these extra bits. In general, they are referred to as error-detecting codes. In specific cases, when the algorithm to create the code is based on addition, they may be called a checksum. We will see that the Internet checksum is appropriately named: It is an error check that uses a summing algorithm. Unfortunately, the word “checksum” is often used imprecisely to mean any form of error-detecting code, including CRCs. This can be confusing, so we urge you to use the word “checksum” only to apply to codes that actually do use addition and to use “error-detecting code” to refer to the general class of codes described in this section. 2.4.1 Two-Dimensional Parity Two-dimensional parity is exactly what the name suggests. It is based on “simple” (one-dimensional) parity, which usually involves adding one extra bit to a 7-bit code to balance the number of 1s in the byte. For example, odd parity sets the eighth bit to 1 if needed to give an odd number of 1s in the byte, and even parity sets the eighth bit to 1 if needed to give an even number of 1s in the byte. Two-dimensional parity does a similar calculation for each bit position across each of the bytes contained in the frame. This results in an extra parity byte for the entire frame, in addition to a parity bit for each byte. Figure 2.16 illustrates how two-dimensional even parity works for an example frame containing 6 bytes of data. Notice that the third bit of the parity byte is 1 since there are an odd number of 1s in the third bit across the 6 bytes in the frame. It can be shown that two-dimensional parity catches all 1-, 2-, and 3-bit errors, and most 4-bit errors. In this case, we have added 14 bits of redundant information 90 2 Direct Link Networks 1011110 1 1101001 0 0101001 1 1011111 0 0110100 1 0001110 1 1111011 0 Parity bits Parity byte Data Figure 2.16 Two-dimensional parity. to a 42-bit message, and yet we have stronger protection against common errors than the “repetition code” described above. 2.4.2 Internet Checksum Algorithm A second approach to error detection is exemplified by the Internet checksum. Although it is not used at the link level, it nevertheless provides the same sort of functionality as CRCs and parity, so we discuss it here. We will see examples of its use in Sections 4.1, 5.1, and 5.2. The idea behind the Internet checksum is very simple—you add up all the words that are transmitted and then transmit the result of that sum. The result is called the checksum. The receiver performs the same calculation on the received data and compares the result with the received checksum. If any transmitted data, including the checksum itself, is corrupted, then the results will not match, so the receiver knows that an error occurred. You can imagine many different variations on the basic idea of a checksum. The exact scheme used by the Internet protocols works as follows. Consider the data being checksummed as a sequence of 16-bit integers. Add them together using 16-bit ones complement arithmetic (explained below) and then take the ones complement of the result. That 16-bit number is the checksum. In ones complement arithmetic, a negative integer −x is represented as the complement of x; that is, each bit of x is inverted. When adding numbers in ones complement arithmetic, a carryout from the most significant bit needs to be added to the result. Consider, for example, the addition of −5 and −3 in ones complement arithmetic on 4-bit integers. +5 is 0101, so −5 is 1010; +3 is 0011, so −3 is 1100. 2.4 Error Detection 91 If we add 1010 and 1100 ignoring the carry, we get 0110. In ones complement arith- metic, the fact that this operation caused a carry from the most significant bit causes us to increment the result, giving 0111, which is the ones complement representation of −8 (obtained by inverting the bits in 1000), as we would expect. The following routine gives a straightforward implementation of the Internet’s checksum algorithm. The count argument gives the length of buf measured in 16-bit units. The routine assumes that buf has already been padded with 0s to a 16-bit boundary. u_short cksum(u_short *buf, int count) { register u_long sum = 0; while (count--) { sum += *buf++; if (sum & 0xFFFF0000) { /* carry occurred, so wrap around */ sum &= 0xFFFF; sum++; } } return ˜ (sum & 0xFFFF); } This code ensures that the calculation uses ones complement arithmetic, rather than the twos complement that is used in most machines. Note the if statement inside the while loop. If there is a carry into the top 16 bits of sum, then we increment sum just as in the previous example. Compared to our repetition code, this algorithm scores well for using a small number of redundant bits—only 16 for a message of any length—but it does not score extremely well for strength of error detection. For example, a pair of single-bit errors, one of which increments a word, one of which decrements another word by the same amount, will go undetected. The reason for using an algorithm like this in spite of its relatively weak protection against errors (compared to a CRC, for example) is simple: This algorithm is much easier to implement in software. Experience in the ARPANET suggested that a checksum of this form was adequate. One reason it is adequate is that this checksum is the last line of defense in an end-to-end protocol; the majority of errors are picked up by stronger error detection algorithms, such as CRCs, at the link level. 92 2 Direct Link Networks 2.4.3 Cyclic Redundancy Check It should be clear by now that a major goal in designing error detection algorithms is to maximize the probability of detecting errors using only a small number of redundant bits. Cyclic redundancy checks use some fairly powerful mathematics to achieve this goal. For example, a 32-bit CRC gives strong protection against common bit errors in messages that are thousands of bytes long. The theoretical foundation of the cyclic redundancy check is rooted in a branch of mathematics called finite fields. While this may sound daunting, the basic ideas can be easily understood. To start, think of an (n + 1)-bit mes- sage as being represented by a polynomial of degree n, that is, a polynomial whose highest-order term is xn. The message is rep- resented by a polynomial by using the value of each bit in the message as the coefficient for each term in the polynomial, starting with the most significant bit to represent the highest-order term. For example, an 8-bit message consisting of the bits 10011010 cor- responds to the polynomial M(x) = 1 × x7 + 0 × x6 + 0 × x5 + 1 × x4 + 1 × x3 + 0 × x2 + 1 × x1 + 0 × x0 = x7 + x4 + x3 + x1 Simple Probability Calculations When dealing with network er- rors and other unlikely (we hope) events, we often have use for simple back-of-the-envelope probability estimates. A useful approximation here is that if two independent events have small probabilities p and q, then the probability of either event is p + q; the exact answer is 1 − (1 − p)(1 − q) = p + q − pq. For p = q = .01, this estimate is .02, while the exact value is .0199. For a simple application of this, suppose that the per-bit error rate on a link is 1 in 107. Assuming bit errors are all independent (which they aren’t), we can estimate that the probability of at least one error in a 10,000-bit packet is 104/107 =10−3. The exact answer, computed as 1 − P(no errors), would be 1 −(1 − 10−7)10,000 = .00099950. For a slightly more complex application, we compute the proba- bility of two errors in such a packet; this is the probability of an error We can thus think of a sender and a receiver as exchanging polynomials with each other. For the purposes o f calculating a CRC, a sender and receiver have to agree on a divisor polynomial, C(x). C(x) is a polynomial of degree k. For example, suppose C(x) = x3 + x2 + 1. In this case, k = 3. The answer to the question “Where did 2.4 Error Detection 93 that would sneak past a 1-parity-bit checksum. Let Eij be the event that bits i and j are bad, for 0 ≤ i < j < 104; the probability of this event is about p = 10−7 × 10−7 = 10−14. For a fixed j, the number of events Eij with i < j is j; adding up the number of these events for all j < 104, we get 1+2+···+(104 −1) ≈1 2 108. The final probability is thus 1 2 108 × 10−14 = 1 2 10−6. Note that had we attempted to estimate P(two errors) = P(first error) × P(second error), and taken these last two to be P(one error) = 10−3, we would have obtained 10−6 here, which is rather far off; the problem with this approach is that not all i are equally likely to be the position of the first error. Or, looked at another way, we have overstated the true probability by a factor of two because we counted errors at positions (i, j) and ( j, i) separately when they should only be counted once. C(x) come from?” is, in most practical cases, “You look it up in a book.” In fact, the choice of C(x) has a significant impact on what types of errors can be reliably detected, as we discuss below. There are a handful of divisor polynomials that are very good choices for various environments, and the exact choice is normally made as part of the protocol design. For example, the Ethernet standard uses a well-known polynomial of degree 32. When a sender wishes to transmit a message M(x) that is n + 1 bits long, what is actually sent is the (n + 1)-bit message plus k bits. We call the complete transmit- ted message, including the redundant bits, P(x). What we are going to do is contrive to make the polynomial representing P(x) exactly divisible by C(x); we explain how this is achieved below. If P(x) is transmit- ted over a link and there are no errors intro- duced during transmission, then the receiver should be able to divide P(x)byC(x) exact- ly, leaving a remainder of zero. On the other hand, if some error is introduced into P(x) during transmission, then in all likelihood the received polynomial will no longer be exactly divisible by C(x), and thus the re- ceiver will obtain a nonzero remainder, im- plying that an error has occurred. It will help to understand the follow- ing if you know a little about polynomial arithmetic; it is just slightly different from normal integer arithmetic. We are dealing with a special class of polynomial arithmetic here, where coefficients may be only one or zero, and operations on the coefficients are performed using modulo 2 arithmetic. This is referred to as “polynomial arithmetic modulo 2.” Since this is a networking book, not a mathematics text, let’s focus on the key properties of this type of arithmetic for our purposes (which we ask you to accept on faith): 94 2 Direct Link Networks ■ Any polynomial B(x) can be divided by a divisor polynomial C(x)ifB(x)is of higher degree than C(x). ■ Any polynomial B(x) can be divided once by a divisor polynomial C(x)if B(x) is of the same degree as C(x). ■ The remainder obtained when B(x) is divided by C(x) is obtained by sub- tracting C(x) from B(x). ■ To subtract C(x) from B(x), we simply perform the exclusive-OR (XOR) operation on each pair of matching coefficients. For example, the polynomial x3 + 1 can be divided by x3 + x2 + 1 (because they are both of degree 3) and the remainder would be 0×x3 +1×x2 +0×x1 +0×x0 = x2 (obtained by XORing the coefficients of each term). In terms of messages, we could say that 1001 can be divided by 1101 and leaves a remainder of 0100. You should be able to see that the remainder is just the bitwise exclusive-OR of the two messages. Now that we know the basic rules for dividing polynomials, we are able to do long division, which is necessary to deal with longer messages. An example appears below. Recall that we wanted to create a polynomial for transmission that is derived from the original message M(x), is k bits longer than M(x), and is exactly divisible by C(x). We can do this in the following way: 1 Multiply M(x)byxk; that is, add k zeroes at the end of the message. Call this zero-extended message T(x). 2 Divide T(x)byC(x) and find the remainder. 3 Subtract the remainder from T(x). It should be obvious that what is left at this point is a message that is exactly divisible by C(x). We may also note that the resulting message consists of M(x) fol- lowed by the remainder obtained in step 2, because when we subtracted the remainder (which can be no more than k bits long), we were just XORing it with the k zeroes added in step 1. This part will become clearer with an example. Consider the message x7 + x4 + x3 + x1, or 10011010. We begin by multiplying by x3, since our divisor polynomial is of degree 3. This gives 10011010000. We divide this by C(x), which corresponds to 1101 in this case. Figure 2.17 shows the polynomial long division operation. Given the rules of polynomial arithmetic described above, the long division operation proceeds much as it would if we were dividing integers. Thus in the first step of our example, we see that the divisor 1101 divides once into the first four bits of the message (1001), since they are of the same degree, and leaves a remainder of 100 (1101 XOR 1001). The next step is to bring down a digit from 2.4 Error Detection 95 Generator 1101 11111001 10011010000 Message 1101 1001 1101 1000 1101 1011 1101 1100 1101 1000 1101 101 Remainder Figure 2.17 CRC calculation using polynomial long division. the message polynomial until we get another polynomial with the same degree as C(x), in this case 1001. We calculate the remainder again (100) and continue until the calculation is complete. Note that the “result” of the long division, which appears at the top of the calculation, is not really of much interest—it is the remainder at the end that matters. You can see from the very bottom of Figure 2.17 that the remainder of the example calculation is 101. So we know that 10011010000 minus 101 would be exactly divisible by C(x), and this is what we send. The minus operation in polynomial arithmetic is the logical XOR operation, so we actually send 10011010101. As noted above, this turns out to be just the original message with the remainder from the long division calculation appended to it. The recipient divides the received polynomial by C(x) and, if the result is 0, concludes that there were no errors. If the result is nonzero, it may be necessary to discard the errored message; with some codes, it may be possible to correct a small error (e.g., if the error affected only one bit). A code that enables error correction is called an error-correcting code (ECC). Now we will consider the question of where the polynomial C(x) comes from. Intuitively, the idea is to select this polynomial so that it is very unlikely to divide evenly into a message that has errors introduced into it. If the transmitted message is P(x), we may think of the introduction of errors as the addition of another polynomial E(x), so the recipient sees P(x) + E(x). The only way that an error could slip by undetected would be if the received message could be evenly divided by C(x), and since we know that P(x) can be evenly divided by C(x), this could only happen if E(x) can be divided evenly by C(x). The trick is to pick C(x) so that this is very unlikely for common types of errors. 96 2 Direct Link Networks CRC C(x) CRC-8 x8 + x2 + x1 + 1 CRC-10 x10 + x9 + x5 + x4 + x1 + 1 CRC-12 x12 + x11 + x3 + x2 + 1 CRC-16 x16 + x15 + x2 + 1 CRC-CCITT x16 + x12 + x5 + 1 CRC-32 x32 + x26 + x23 + x22 + x16 + x12 + x11 + x10 + x8 + x7 + x5 + x4 + x2 + x + 1 Table 2.5 Common CRC polynomials. One common type of error is a single-bit error, which can be expressed as E(x) =xi when it affects bit position i. If we select C(x) such that the first and the last term are nonzero, then we already have a two-term polynomial that cannot divide evenly into the one term E(x). Such a C(x) can, therefore, detect all single-bit errors. In general, it is possible to prove that the following types of errors can be detected by a C(x) with the stated properties: ■ All single-bit errors, as long as the xk and x0 terms have nonzero coef- ficients. ■ All double-bit errors, as long as C(x) has a factor with at least three terms. ■ Any odd number of errors, as long as C(x) contains the factor (x + 1). ■ Any “burst” error (i.e., sequence of consecutive errored bits) for which the length of the burst is less than k bits. (Most burst errors of larger than k bits can also be detected.) Six versions of C(x) are widely used in link-level protocols (shown in Table 2.5). Error Detection or Error Correction? We have mentioned that it is possi- ble to use codes that not only detect the presence of errors but also en- able errors to be corrected. Since the details of such codes require yet more complex mathematics than that required to understand CRCs, we will not dwell on them here. However, it is worth considering the merits of correction versus de- tection. At first glance, it would seem that correction is always better, since with detection we are forced to throw away the message and, 2.5 Reliable Transmission 97 x0 x1 XOR gate x2 Message Figure 2.18 CRC calculation using shift register. For example, the Ethernet and 802.5 networks described later in this chapter use CRC-32, while HDLC uses CRC-CCITT. ATM, as described in Chapter 3, uses CRC- 8, CRC-10, and CRC-32. Finally, we note that the CRC algorithm, while seemingly complex, is easily implemented in hardware using a k-bit shift register and XOR gates. The number of bits in the shift register equals the degree of the generator polynomial (k). Figure 2.18 shows the hardware that would be used for the generator x3 +x2 +1 from our previous example. The message is shifted in from the left, beginning with the most significant bit and ending with the string of k zeroes that is attached to the message, just as in general, ask for another copy to be transmitted. This uses up band- width and may introduce latency while waiting for the retransmis- sion. However, there is a downside to correction: It generally requires a greater number of redundant bits to send an error-correcting code that is as strong (that is, able to cope with the same range of errors) as a code that only detects errors. Thus, while error detection requires more bits to be sent when errors occur, error correction requires more bits to be sent all the time. As a result, in the long division example. When all the bits have been shifted in and appropriately XORed, the register contains the remainder, that is, the CRC (most significant bit on the right). The position of the XOR gates is de- termined as follows: If the bits in the shift register are labelled 0 through k − 1, left to right, then put an XOR gate in front of bit n if there is a term xn in the generator polynomial. Thus, we see an XOR gate in front of positions 0 and 2 for the generator x3 + x2 + x0. 2.5 Reliable Transmission As we saw in the previous section, frames are sometimes corrupted while in tran- sit, with an error code like CRC used to detect such errors. While some error codes are strong enough also to correct errors, in practice the overhead is typically too 98 2 Direct Link Networks large to handle the range of bit and burst errors that can be introduced on a network link. Even when error-correcting codes are used (e.g., on wireless links), some errors will be too severe to be corrected. As a result, some corrupt frames must be discarded. A link-level protocol that wants to deliver frames reliably must somehow recover from these discarded (lost) frames. This is usually accomplished using a combination of two fundamental mechanisms—acknowledgments and timeouts. An acknowledgment (ACK for short) is a small control frame that a protocol sends back to its peer saying that it has received an earlier frame. By control frame we mean a header without any data, although a protocol can piggyback an ACK on a data frame it just happens to be sending in the opposite direction. The receipt of an acknowledgment indicates to the sender of the original frame that its frame was successfully delivered. If the sender does not receive an acknowledgment after a reasonable amount of time, then it retransmits the original frame. This action of waiting a reasonable amount of time is called a timeout. The general strategy of using acknowledgments and timeouts to implement re- liable delivery is sometimes called automatic repeat request (normally abbreviated ARQ). This section describes three different ARQ algorithms using generic language; that is, we do not give detailed information about a particular protocol’s header fields. 2.5.1 Stop-and-Wait The simplest ARQ scheme is the stop-and- wait algorithm. The idea of stop-and-wait is straightforward: After transmitting one frame, the sender waits for an acknowledg- ment before transmitting the next frame. If the acknowledgment does not arrive after a certain period of time, the sender times out and retransmits the original frame. Figure 2.19 illustrates four different scenarios that result from this basic algo- rithm. This figure is a timeline, a com- mon way to depict a protocol’s behavior. The sending side is represented on the left, the receiving side is depicted on the right, and time flows from top to bottom. Fig- ure 2.19(a) shows the situation in which the ACK is received before the timer ex- pires, (b) and (c) show the situation in which error correction tends to be most useful when (1) errors are quite probable, as they may be, for ex- ample, in a wireless environment, or (2) the cost of retransmission is too high, for example, because of the latency involved in retransmit- ting a packet over a satellite link. The use of error-correcting codes in networking is sometimes referred to as forward error cor- rection (FEC) because the cor- rection of errors is handled “in advance” by sending extra infor- mation, rather than waiting for errors to happen and dealing with them later by retransmission. 2.5 Reliable Transmission 99 the original frame and the ACK, respectively, are lost, and (d) shows the situation in which the timeout fires too soon. Recall that by “lost” we mean that the frame was corrupted while in transit, that this corruption was detected by an error code on the receiver, and that the frame was subsequently discarded. There is one important subtlety in the stop-and-wait algorithm. Suppose the sender sends a frame and the receiver acknowledges it, but the acknowledgment is either lost or delayed in arriving. This situation is illustrated in timelines (c) and (d) of Figure 2.19. In both cases, the sender times out and retransmits the original frame, Sender Receiver Frame ACK TimeoutTime Sender Receiver Frame ACK Timeout Frame ACKTimeout Sender Receiver Frame ACKTimeout Frame ACKTimeout Sender Receiver Frame Timeout Frame ACKTimeout (a) (c) (b) (d) Figure 2.19 Timeline showing four different scenarios for the stop-and-wait algorithm. (a) The ACK is received before the timer expires; (b) the original frame is lost; (c) the ACK is lost; (d) the timeout fires too soon. 100 2 Direct Link Networks Sender Receiver Frame 0 ACK 0 Time Frame 1 ACK 1 Frame 0 ACK 0 … Figure 2.20 Timeline for stop-and-wait with 1-bit sequence number. but the receiver will think that it is the next frame, since it correctly received and acknowledged the first frame. This has the potential to cause duplicate copies of a frame to be delivered. To address this problem, the header for a stop-and-wait protocol usually includes a 1-bit sequence number—that is, the sequence number can take on the values 0 and 1—and the sequence numbers used for each frame alternate, as illustrated in Figure 2.20. Thus, when the sender retransmits frame 0, the receiver can determine that it is seeing a second copy of frame 0 rather than the first copy of frame 1 and therefore can ignore it (the receiver still acknowledges it, in case the first ACK was lost). The main shortcoming of the stop-and-wait algorithm is that it allows the sender to have only one outstanding frame on the link at a time, and this may be far below the link’s capacity. Consider, for example, a 1.5-Mbps link with a 45-ms round-trip time. This link has a delay × bandwidth product of 67.5 Kb, or approximately 8 KB. Since the sender can send only one frame per RTT, and assuming a frame size of 1 KB, this implies a maximum sending rate of BitsPerFrame ÷ TimePerFrame = 1024 × 8 ÷ 0.045 = 182 Kbps or about one-eighth of the link’s capacity. To use the link fully, then, we’d like the sender to be able to transmit up to eight frames before having to wait for an acknowledgment. 2.5 Reliable Transmission 101 ◮ The significance of the bandwidth × delay product is that it represents the amount of data that could be in transit. We would like to be able to send this much data without waiting for the first acknowledgment. The principle at work here is often referred to as keeping the pipe full. The algorithms presented in the following two subsections do exactly this. 2.5.2 Sliding Window Consider again the scenario in which the link has a delay × bandwidth product of 8 KB and frames are of 1-KB size. We would like the sender to be ready to transmit the ninth frame at pretty much the same moment that the ACK for the first frame arrives. The algorithm that allows us to do this is called sliding window, and an illustrative timeline is given in Figure 2.21. The Sliding Window Algorithm The sliding window algorithm works as follows. First, the sender assigns a sequence number, denoted SeqNum, to each frame. For now, let’s ignore the fact that SeqNum is implemented by a finite-size header field and instead assume that it can grow infinitely large. The sender maintains three variables: The send window size, denoted SWS, gives the upper bound on the number of outstanding (unacknowledged) frames that the sender can transmit; LAR denotes the sequence number of the last acknowledgment received; and LFS denotes the sequence number of the last frame sent. The sender also maintains the following invariant: LFS − LAR ≤ SWS This situation is illustrated in Figure 2.22. Sender Receiver Time …… Figure 2.21 Timeline for the sliding window algorithm. 102 2 Direct Link Networks ≤ SWS LAR LFS …… Figure 2.22 Sliding window on sender. ≤ RWS LFR LAF …… Figure 2.23 Sliding window on receiver. When an acknowledgment arrives, the sender moves LAR to the right, thereby allowing the sender to transmit another frame. Also, the sender associates a timer with each frame it transmits, and it retransmits the frame should the timer expire before an ACK is received. Notice that the sender has to be willing to buffer up to SWS frames since it must be prepared to retransmit them until they are acknowledged. The receiver maintains the following three variables: The receive window size, denoted RWS, gives the upper bound on the number of out-of-order frames that the receiver is willing to accept; LAF denotes the sequence number of the largest acceptable frame; and LFR denotes the sequence number of the last frame received. The receiver also maintains the following invariant: LAF − LFR ≤ RWS This situation is illustrated in Figure 2.23. When a frame with sequence number SeqNum arrives, the receiver takes the following action. If SeqNum ≤ LFR or SeqNum > LAF, then the frame is outside the receiver’s window and it is discarded. If LFR < SeqNum ≤ LAF, then the frame is within the receiver’s window and it is accepted. Now the receiver needs to decide whether or not to send an ACK. Let SeqNumToAck denote the largest sequence number not yet acknowledged, such that all frames with sequence numbers less than or equal to SeqNumToAck have been received. The receiver acknowledges the receipt of SeqNumToAck, even if higher-numbered packets have been received. 2.5 Reliable Transmission 103 This acknowledgment is said to be cumulative. It then sets LFR = SeqNumToAck and adjusts LAF = LFR + RWS. For example, suppose LFR = 5 (i.e., the last ACK the receiver sent was for sequence number 5), and RWS = 4. This implies that LAF = 9. Should frames 7 and 8 arrive, they will be buffered because they are within the receiver’s window. However, no ACK needs to be sent since frame 6 is yet to arrive. Frames 7 and 8 are said to have arrived out of order. (Technically, the receiver could resend an ACK for frame 5 when frames 7 and 8 arrive.) Should frame 6 then arrive—perhaps it is late because it was lost the first time and had to be retransmitted, or perhaps it was simply delayed—the receiver acknowledges frame 8, bumps LFR to 8, and sets LAF to 12. If frame 6 was in fact lost, then a timeout will have occurred at the sender, causing it to retransmit frame 6. We observe that when a timeout occurs, the amount of data in transit decreases, since the sender is unable to advance its window until frame 6 is acknowledged. This means that when packet losses occur, this scheme is no longer keeping the pipe full. The longer it takes to notice that a packet loss has occurred, the more severe this problem becomes. Notice that in this example, the receiver could have sent a negative acknowl- edgment (NAK) for frame 6 as soon as frame 7 arrived. However, this is unnecessary since the sender’s timeout mechanism is sufficient to catch this situation, and sending NAKs adds additional complexity to the receiver. Also, as we mentioned, it would have been legitimate to send additional acknowledgments of frame 5 when frames 7 and 8 arrived; in some cases, a sender can use duplicate ACKs as a clue that a frame was lost. Both approaches help to improve performance by allowing early detection of packet losses. Yet another variation on this scheme would be to use selective acknowledgments. That is, the receiver could acknowledge exactly those frames it has received, rather than just the highest-numbered frame received in order. So, in the above example, the receiver could acknowledge the receipt of frames 7 and 8. Giving more information to the sender makes it potentially easier for the sender to keep the pipe full, but adds complexity to the implementation. The sending window size is selected according to how many frames we want to have outstanding on the link at a given time; SWS is easy to compute for a given delay ×bandwidth product.1 On the other hand, the receiver can set RWS to whatever it wants. Two common settings are RWS = 1, which implies that the receiver will not buffer any frames that arrive out of order, and RWS = SWS, which implies that the receiver can 1Easy, that is, if we know the delay and the bandwidth. Sometimes we do not, and estimating them well is a challenge to protocol designers. We discuss this further in Chapter 5. 104 2 Direct Link Networks buffer any of the frames the sender transmits. It makes no sense to set RWS > SWS since it’s impossible for more than SWS frames to arrive out of order. Finite Sequence Numbers and Sliding Window We now return to the one simplification we introduced into the algorithm—our as- sumption that sequence numbers can grow infinitely large. In practice, of course, a frame’s sequence number is specified in a header field of some finite size. For example, a 3-bit field means that there are eight possible sequence numbers, 0 ...7. This makes it necessary to reuse sequence numbers or, stated another way, sequence numbers wrap around. This introduces the problem of being able to distinguish between different in- carnations of the same sequence numbers, which implies that the number of possible sequence numbers must be larger than the number of outstanding frames allowed. For example, stop-and-wait allowed one outstanding frame at a time and had two distinct sequence numbers. Suppose we have one more number in our space of sequence numbers than we have potentially outstanding frames; that is, SWS ≤ MaxSeqNum − 1, where MaxSeqNum is the number of available sequence numbers. Is this sufficient? The answer depends on RWR.IfRWS = 1, then MaxSeqNum ≥ SWS + 1 is sufficient. If RWS is equal to SWS, then having a MaxSeqNum just one greater than the sending window size is not good enough. To see this, consider the situation in which we have the eight sequence numbers 0 through 7, and SWS = RWS = 7. Suppose the sender transmits frames 0..6, they are successfully received, but the ACKs are lost. The receiver is now expecting frames 7, 0..5, but the sender times out and sends frames 0..6. Unfortunately, the receiver is expecting the second incarnation of frames 0..5, but gets the first incarnation of these frames. This is exactly the situation we wanted to avoid. It turns out that the sending window size can be no more than half as big as the number of available sequence numbers when RWS = SWS, or stated more precisely, SWS < (MaxSeqNum + 1)/2 Intuitively, what this is saying is that the sliding window protocol alternates between the two halves of the sequence number space, just as stop-and-wait alternates between sequence numbers 0 and 1. The only difference is that it continually slides between the two halves rather than discretely alternating between them. Note that this rule is specific to the situation where RWS = SWS. We leave it as an exercise to determine the more general rule that works for arbitrary val- ues of RWS and SWS. Also note that the relationship between the window size and the sequence number space depends on an assumption that is so obvious that it 2.5 Reliable Transmission 105 is easy to overlook, namely, that frames are not reordered in transit. This cannot happen on a direct point-to-point link since there is no way for one frame to overtake another during transmission. However, we will see the sliding window algorithm used in a different environment in Chapter 5, and we will need to devise another rule. Implementation of Sliding Window The following routines illustrate how we might implement the sending and receiving sides of the sliding window algorithm. The routines are taken from a working protocol named, appropriately enough, sliding window protocol (SWP). So as not to concern ourselves with the adjacent protocols in the protocol graph, we denote the protocol sitting above SWP as HLP (high-level protocol) and the protocol sitting below SWP as LINK (link-level protocol). We start by defining a pair of data structures. First, the frame header is very simple: It contains a sequence number (SeqNum) and an acknowledgment number (AckNum). It also contains a Flags field that indicates whether the frame is an ACK or carries data. typedef u_char SwpSeqno; typedef struct { SwpSeqno SeqNum; /* sequence number of this frame */ SwpSeqno AckNum; /* ack of received frame */ u_char Flags; /* up to 8 bits' worth of flags */ } SwpHdr; Next, the state of the sliding window algorithm has the following structure. For the sending side of the protocol, this state includes variables LAR and LFS, as described earlier in this section, as well as a queue that holds frames that have been transmit- ted but not yet acknowledged (sendQ). The sending state also includes a counting semaphore called sendWindowNotFull. We will see how this is used below, but gener- ally a semaphore is a synchronization primitive that supports semWait and semSig- nal operations. Every invocation of semSignal increments the semaphore by 1, and every invocation of semWait decrements s by 1, with the calling process blocked (sus- pended) should decrementing the semaphore cause its value to become less than 0. A process that is blocked during its call to semWait will be allowed to resume as soon as enough semSignal operations have been performed to raise the value of the semaphore above 0. For the receiving side of the protocol, the state includes the variable NFE. This is the next frame expected—the frame with a sequence number one more than the last frame received (LFR), described earlier in this section. There is also a queue that holds frames that have been received out of order (recvQ). Finally, although not shown, 106 2 Direct Link Networks the sender and receiver sliding window sizes are defined by constants SWS and RWS, respectively. typedef struct { /* sender side state: */ SwpSeqno LAR; /* seqno of last ACK received */ SwpSeqno LFS; /* last frame sent */ Semaphore sendWindowNotFull; SwpHdr hdr; /* preinitialized header */ struct sendQ_slot { Event timeout; /* event associated with send-timeout */ Msg msg; } sendQ[SWS]; /* receiver side state: */ SwpSeqno NFE; /* seqno of next frame expected */ struct recvQ_slot { int received; /* is msg valid? */ Msg msg; } recvQ[RWS]; } SwpState; The sending side of SWP is implemented by procedure sendSWP. This routine is rather simple. First, semWait causes this process to block on a semaphore until it is OK to send another frame. Once allowed to proceed, sendSWP sets the sequence number in the frame’s header, saves a copy of the frame in the transmit queue (sendQ), schedules a timeout event to handle the case in which the frame is not acknowledged, and sends the frame to the next-lower-level protocol, which we denote as LINK. One detail worth noting is the call to store swp hdr just before the call to msgAddHdr. This routine translates the C structure that holds the SWP header (state-> hdr) into a byte string that can be safely attached to the front of the message (hbuf). This routine (not shown) must translate each integer field in the header into network byte order and remove any padding that the compiler has added to the C structure. The issue of byte order is discussed more fully in Section 7.1, but for now it is enough to assume that this routine places the most significant bit of a multiword integer in the byte with the highest address. Also, we assume an abstract data type, denoted Msg, that holds a message. The Msg type supports operations like msgAddHdr and msgSaveCopy. Another piece of complexity in this routine is the use of semWait and the send- WindowNotFull semaphore. sendWindowNotFull is initialized to the size of the sender’s sliding window, SWS (this initialization is not shown). Each time the sender trans- mits a frame, the semWait operation decrements this count and blocks the sender should the count go to 0. Each time an ACK is received, the semSignal operation 2.5 Reliable Transmission 107 invoked in deliverSWP (see below) increments this count, thus unblocking any wait- ing sender. static int sendSWP(SwpState *state, Msg *frame) { struct sendQ_slot *slot; hbuf[HLEN]; /* wait for send window to open */ semWait(&state->sendWindowNotFull); state->hdr.SeqNum = ++state->LFS; slot = &state->sendQ[state->hdr.SeqNum % SWS]; store_swp_hdr(state->hdr, hbuf); msgAddHdr(frame, hbuf, HLEN); msgSaveCopy(&slot->msg, frame); slot->timeout = evSchedule(swpTimeout, slot, SWP_SEND_TIMEOUT); return sendLINK(frame); } Now to SWP’s protocol-specific implementation of the deliver operation, which is given in procedure deliverSWP. This routine actually handles two different kinds of incoming messages: ACKs for frames sent earlier from this node and data frames arriving at this node. In a sense, the ACK half of this routine is the counterpart to the sender side of the algorithm given in sendSWP. A decision as to whether the incoming message is an ACK or a data frame is made by checking the Flags field in the header. Note that this particular implementation does not support piggybacking ACKs on data frames. When the incoming frame is an ACK, deliverSWP simply finds the slot in the transmit queue (sendQ) that corresponds to the ACK, cancels the timeout event, and frees the frame saved in that slot. This work is actually done in a loop since the ACK may be cumulative. The only other thing to notice about this case is the call to subroutine swpInWindow. This subroutine, which is given below, ensures that the sequence number for the frame being acknowledged is within the range of ACKs that the sender currently expects to receive. When the incoming frame contains data, deliverSWP first calls msgStripHdr and load swp hdr to extract the header from the frame. Routine load swp hdr is the coun- terpart to store swp hdr discussed earlier; it translates a byte string into the C data structure that holds the SWP header. deliverSWP then calls swpInWindow to make sure the sequence number of the frame is within the range of sequence numbers that it expects. If it is, the routine loops over the set of consecutive frames it has received and passes them up to the higher-level protocol by invoking the deliverHLP routine. It also 108 2 Direct Link Networks sends a cumulative ACK back to the sender, but does so by looping over the receive queue (it does not use the SeqNumToAck variable used in the prose description given earlier in this section). static int deliverSWP(SwpState state, Msg *frame) { SwpHdr hdr; char *hbuf; hbuf = msgStripHdr(frame, HLEN); load_swp_hdr(&hdr, hbuf) if (hdr->Flags & FLAG_ACK_VALID) { /* received an acknowledgment---do SENDER side */ if (swpInWindow(hdr.AckNum, state->LAR + 1, state->LFS)) { do { struct sendQ_slot *slot; slot = &state->sendQ[++state->LAR % SWS]; evCancel(slot->timeout); msgDestroy(&slot->msg); semSignal(&state->sendWindowNotFull); } while (state->LAR != hdr.AckNum); } } if (hdr.Flags & FLAG_HAS_DATA) { struct recvQ_slot *slot; /* received data packet---do RECEIVER side */ slot = &state->recvQ[hdr.SeqNum % RWS]; if (!swpInWindow(hdr.SeqNum, state->NFE, state->NFE + RWS - 1)) { /* drop the message */ return SUCCESS; } msgSaveCopy(&slot->msg, frame); slot->received = TRUE; if (hdr.SeqNum == state->NFE) { 2.5 Reliable Transmission 109 Msg m; while (slot->received) { deliverHLP(&slot->msg); msgDestroy(&slot->msg); slot->received = FALSE; slot = &state->recvQ[++state->NFE % RWS]; } /* send ACK: */ prepare_ack(&m, state->NFE - 1); sendLINK(&m); msgDestroy(&m); } } return SUCCESS; } Finally, swpInWindow is a simple subroutine that checks to see if a given sequence number falls between some minimum and maximum sequence number. static bool swpInWindow(SwpSeqno seqno, SwpSeqno min, SwpSeqno max) { SwpSeqno pos, maxpos; pos = seqno - min; /* pos *should* be in range [0..MAX)*/ maxpos = max - min + 1; /* maxpos is in range [0..MAX]*/ return pos < maxpos; } Frame Order and Flow Control The sliding window protocol is perhaps the best-known algorithm in computer net- working. What is easily confusing about the algorithm, however, is that it can be used to serve three different roles. The first role is the one we have been concentrating on in this section—to reliably deliver frames across an unreliable link. (In general, the algorithm can be used to reliably deliver messages across an unreliable network.) This is the core function of the algorithm. The second role that the sliding window algorithm can serve is to preserve the order in which frames are transmitted. This is easy to do at the receiver—since each frame has a sequence number, the receiver just makes sure that it does not pass a frame up to the next-higher-level protocol until it has already passed up all frames with a smaller sequence number. That is, the receiver buffers (i.e., does not pass along) out-of-order frames. The version of the sliding window algorithm described in this 110 2 Direct Link Networks section does preserve frame order, although we could imagine a variation in which the receiver passes frames to the next protocol without waiting for all earlier frames to be delivered. A question we should ask ourselves is whether we really need the sliding window protocol to keep the frames in order, or whether, instead, this is un- necessary functionality at the link level. Unfortunately, we have not yet seen enough of the network architecture to answer this question; we first need to understand how a sequence of point-to-point links is connected by switches to form an end-to-end path. The third role that the sliding window algorithm sometimes plays is to support flow control—a feedback mechanism by which the receiver is able to throttle the sender. Such a mechanism is used to keep the sender from overrunning the receiver, that is, from transmitting more data than the receiver is able to process. This is usually accomplished by augmenting the sliding window protocol so that the receiver not only acknowledges frames it has received, but also informs the sender of how many frames it has room to receive. The number of frames that the receiver is capable of receiving corresponds to how much free buffer space it has. As in the case of ordered delivery, we need to make sure that flow control is necessary at the link level before incorporating it into the sliding window protocol.◮ One important concept to take away from this discussion is the system design principle we call separation of concerns. That is, you must be careful to distinguish between different functions that are sometimes rolled together in one mechanism, and you must make sure that each function is necessary and being supported in the most effective way. In this particular case, reliable delivery, ordered delivery, and flow control are sometimes combined in a single sliding window protocol, and we should ask ourselves if this is the right thing to do at the link level. With this question in mind, we revisit the sliding window algorithm in Chapter 3 (we show how X.25 networks use it to implement hop-by-hop flow control) and in Chapter 5 (we describe how TCP uses it to implement a reliable byte-stream channel). 2.5.3 Concurrent Logical Channels The data link protocol used in the ARPANET provides an interesting alternative to the sliding window protocol, in that it is able to keep the pipe full while still using the simple stop-and-wait algorithm. One important consequence of this approach is that the frames sent over a given link are not kept in any particular order. The protocol also implies nothing about flow control. The idea underlying the ARPANET protocol, which we refer to as concurrent logical channels, is to multiplex several logical channels onto a single point-to-point link and to run the stop-and-wait algorithm on each of these logical channels. There is no relationship maintained among the frames sent on any of the logical channels, yet 2.6 Ethernet (802.3) 111 because a different frame can be outstanding on each of the several logical channels, the sender can keep the link full. More precisely, the sender keeps 3 bits of state for each channel: a boolean, saying whether the channel is currently busy; the 1-bit sequence number to use the next time a frame is sent on this logical channel; and the next sequence number to expect on a frame that arrives on this channel. When the node has a frame to send, it uses the lowest idle channel, and otherwise it behaves just like stop-and-wait. In practice, the ARPANET supported 8 logical channels over each ground link and 16 over each satellite link. In the ground-link case, the header for each frame included a 3-bit channel number and a 1-bit sequence number, for a total of 4 bits. This is exactly the number of bits the sliding window protocol requires to support up to eight outstanding frames on the link when RWS = SWS. 2.6 Ethernet (802.3) The Ethernet is easily the most successful local area networking technology of the last 20 years. Developed in the mid-1970s by researchers at the Xerox Palo Alto Research Center (PARC), the Ethernet is a working example of the more general Carrier Sense Multiple Access with Collision Detect (CSMA/CD) local area network technology. As indicated by the CSMA name, the Ethernet is a multiple-access network, meaning that a set of nodes send and receive frames over a shared link. You can, therefore, think of an Ethernet as being like a bus that has multiple stations plugged into it. The “carrier sense” in CSMA/CD means that all the nodes can distinguish between an idle and a busy link, and “collision detect” means that a node listens as it transmits and can therefore detect when a frame it is transmitting has interfered (collided) with a frame transmitted by another node. The Ethernet has its roots in an early packet radio network, called Aloha, de- veloped at the University of Hawaii to support computer communication across the Hawaiian Islands. Like the Aloha network, the fundamental problem faced by the Ethernet is how to mediate access to a shared medium fairly and efficiently (in Aloha the medium was the atmosphere, while in Ethernet the medium is a coax cable). That is, the core idea in both Aloha and the Ethernet is an algorithm that controls when each node can transmit. Digital Equipment Corporation and Intel Corporation joined Xerox to define a 10-Mbps Ethernet standard in 1978. This standard then formed the basis for IEEE standard 802.3. With one exception that we will see in Section 2.6.2, it is fair to view the 1978 Ethernet standard as a proper subset of the 802.3 standard; 802.3 additionally defines a much wider collection of physical media over which Ethernet can operate, and more recently, it has been extended to include a 100-Mbps version called Fast 112 2 Direct Link Networks Transceiver Ethernet cable Adaptor Host Figure 2.24 Ethernet transceiver and adaptor. Ethernet and a 1000-Mbps version called Gigabit Ethernet. The rest of this section focuses on 10-Mbps Ethernet, since it is typically used in multiple-access mode and we are interested in how multiple hosts share a single link. Both 100-Mbps and 1000- Mbps Ethernets are designed to be used in full-duplex, point-to-point configurations, which means that they are typically used in switched networks, as described in the next chapter. 2.6.1 Physical Properties An Ethernet segment is implemented on a coaxial cable of up to 500 m. This cable is similar to the type used for cable TV, except that it typically has an impedance of 50 ohms instead of cable TV’s 75 ohms. Hosts connect to an Ethernet segment by tapping into it; taps must be at least 2.5 m apart. A transceiver—a small device directly attached to the tap—detects when the line is idle and drives the signal when the host is transmitting. It also receives incoming signals. The transceiver is, in turn, connected to an Ethernet adaptor, which is plugged into the host. All the logic that makes up the Ethernet protocol, as described in this section, is implemented in the adaptor (not the transceiver). This configuration is shown in Figure 2.24. Multiple Ethernet segments can be joined together by repeaters. A repeater is a de- vice that forwards digital signals, much like an amplifier forwards analog signals. How- ever,no more than four repeaters may be positioned between any pair of hosts, meaning that an Ethernet has a total reach of only 2500 m. For example, using just two re- peaters between any pair of hosts supports a configuration similar to the one illustrated in Figure 2.25, that is, a segment running down the spine of a building with a segment on each floor. All told, an Ethernet is limited to supporting a maximum of 1024 hosts. 2.6 Ethernet (802.3) 113 Repeater Host … … … … Figure 2.25 Ethernet repeater. Any signal placed on the Ethernet by a host is broadcast over the entire network; that is, the signal is propagated in both directions, and repeaters forward the signal on all outgoing segments. Terminators attached to the end of each segment absorb the signal and keep it from bouncing back and interfering with trailing signals. The Ethernet uses the Manchester encoding scheme described in Section 2.2. In addition to the system of segments and repeaters just described, alternative technologies have been introduced over the years. For example, rather than using a 50-ohm coax cable, an Ethernet can be constructed from a thinner cable known as 10Base2; the original cable is called 10Base5 (the two cables are commonly called thin- net and thick-net, respectively). The “10” in 10Base2 means that the network operates at 10 Mbps, “Base” refers to the fact that the cable is used in a baseband system, and the “2” means that a given segment can be no longer than 200 m (a segment of the original 10Base5 cable can be up to 500 m long). Today, a third cable technology is predominantly used, called 10BaseT, where the “T” stands for twisted pair. Typically, Category 5 twisted pair wiring is used. A 10BaseT segment is usually limited to under 100 m in length. (Both 100-Mbps and 1000-Mbps Ethernets also run over Category 5 twisted pair, up to distances of 100 m.) 114 2 Direct Link Networks Hub Hub Figure 2.26 Ethernet hub. Because the cable is so thin, you do not tap into a 10Base2 or 10BaseT cable in the same way as you would with 10Base5 cable. With 10Base2, a T-joint is spliced into the cable. In effect, 10Base2 is used to daisy-chain a set of hosts together. With 10BaseT, the common configuration is to have several point-to-point segments coming out of a multiway repeater, sometimes called a hub, as illustrated in Figure 2.26. Multiple 100-Mbps Ethernet segments can also be connected by a hub, but the same is not true of 1000-Mbps segments. It is important to understand that whether a given Ethernet spans a single seg- ment, a linear sequence of segments connected by repeaters, or multiple segments connected in a star configuration by a hub, data transmitted by any one host on that Ethernet reaches all the other hosts. This is the good news. The bad news is that all these hosts are competing for access to the same link, and as a consequence, they are said to be in the same collision domain. 2.6.2 Access Protocol We now turn our attention to the algorithm that controls access to the shared Ethernet link. This algorithm is commonly called the Ethernet’s media access control (MAC). It is typically implemented in hardware on the network adaptor. We will not describe the hardware per se, but instead focus on the algorithm it implements. First, however, we describe the Ethernet’s frame format and addresses. Frame Format Each Ethernet frame is defined by the format given in Figure 2.27. The 64-bit preamble allows the receiver to synchronize with the signal; it is a sequence of alternating 0s and 1s. Both the source and destination hosts are identified with a 48-bit address. The packet type field serves as the demultiplexing key; that is, it identifies to which of possibly many higher-level protocols this frame should be delivered. Each frame contains up to 1500 bytes of data. Minimally, a frame must contain at least 46 bytes of data, even if this means the host has to pad the frame before transmitting it. The reason 2.6 Ethernet (802.3) 115 Dest addr 64 48 32 CRCPreamble Src addr Type Body 1648 Figure 2.27 Ethernet frame format. for this minimum frame size is that the frame must be long enough to detect a collision; we discuss this more below. Finally, each frame includes a 32-bit CRC. Like the HDLC protocol described in Section 2.3.2, the Ethernet is a bit-oriented framing protocol. Note that from the host’s perspective, an Ethernet frame has a 14-byte header: two 6-byte addresses and a 2-byte type field. The sending adaptor attaches the preamble, CRC, and postamble before transmitting, and the receiving adaptor removes them. The frame format just described is taken from the Digital-Intel-Xerox Ethernet standard. The 802.3 frame format is exactly the same, except it substitutes a 16-bit length field for the 16-bit type field. 802.3 is usually paired with an encapsulation standard that defines a type field used to demultiplex incoming frames. This type field is the first thing in the data portion of the 802.3 frames; that is, it immediately follows the 802.3 header. Fortunately, since the Ethernet standard has avoided using any type values less than 1500 (the maximum length found in an 802.3 header), and the type and length fields are in the same location in the header, it is possible for a single device to accept both formats, and for the device driver running on the host to interpret the last 16 bits of the header as either a type or a length. In practice, most hosts follow the Digital-Intel-Xerox format and interpret this field as the frame’s type. Addresses Each host on an Ethernet—in fact, every Ethernet host in the world—has a unique Ethernet address. Technically, the address belongs to the adaptor, not the host; it is usually burned into ROM. Ethernet addresses are typically printed in a form humans can read as a sequence of six numbers separated by colons. Each number corresponds to 1 byte of the 6-byte address and is given by a pair of hexadecimal digits, one for each of the 4-bit nibbles in the byte; leading 0s are dropped. For example, 8:0:2b:e4:b1:2 is the human-readable representation of Ethernet address 00001000 00000000 00101011 11100100 10110001 00000010 To ensure that every adaptor gets a unique address, each manufacturer of Eth- ernet devices is allocated a different prefix that must be prepended to the address on every adaptor they build. For example, Advanced Micro Devices has been assigned the 116 2 Direct Link Networks 24-bit prefix x080020 (or 8:0:20). A given manufacturer then makes sure the address suffixes it produces are unique. Each frame transmitted on an Ethernet is received by every adaptor connected to that Ethernet. Each adaptor recognizes those frames addressed to its address and passes only those frames on to the host. (An adaptor can also be programmed to run in promiscuous mode, in which case it delivers all received frames to the host, but this is not the normal mode.) In addition to these unicast addresses, an Ethernet address consisting of all 1s is treated as a broadcast address; all adaptors pass frames addressed to the broadcast address up to the host. Similarly, an address that has the first bit set to 1 but is not the broadcast address is called a multicast address. A given host can program its adaptor to accept some set of multicast addresses. Multicast addresses are used to send messages to some subset of the hosts on an Ethernet (e.g., all file servers). To summarize, an Ethernet adaptor receives all frames and accepts ■ frames addressed to its own address ■ frames addressed to the broadcast address ■ frames addressed to a multicast address, if it has been instructed to listen to that address ■ all frames, if it has been placed in promiscuous mode It passes to the host only the frames that it accepts. Transmitter Algorithm As we have just seen, the receiver side of the Ethernet protocol is simple; the real smarts are implemented at the sender’s side. The transmitter algorithm is defined as follows. When the adaptor has a frame to send and the line is idle, it transmits the frame immediately; there is no negotiation with the other adaptors. The upper bound of 1500 bytes in the message means that the adaptor can occupy the line for only a fixed length of time. When an adaptor has a frame to send and the line is busy, it waits for the line to go idle and then transmits immediately.2 The Ethernet is said to be a 1-persistent protocol because an adaptor with a frame to send transmits with probability 1 whenever a busy line goes idle. In general, a p-persistent algorithm transmits with probability 0 ≤ p ≤ 1 after a line becomes idle, and defers with probability q = 1 − p. The reasoning behind choosing a p < 1 is that there might be multiple adaptors waiting for the busy line 2To be more precise, all adaptors wait 9.6 μs after the end of one frame before beginning to transmit the next frame. This is true for the sender of the first frame, as well as those nodes listening for the line to become idle. 2.6 Ethernet (802.3) 117 to become idle, and we don’t want all of them to begin transmitting at the same time. If each adaptor transmits immediately with a probability of, say, 33%, then up to three adaptors can be waiting to transmit and the odds are that only one will begin transmitting when the line becomes idle. Despite this reasoning, an Ethernet adaptor always transmits immediately after noticing that the network has become idle and has been very effective in doing so. To complete the story about p-persistent protocols for the case when p < 1, you might wonder how long a sender that loses the coin flip (i.e., decides to defer) has to wait before it can transmit. The answer for the Aloha network, which originally developed this style of protocol, was to divide time into discrete slots, with each slot corresponding to the length of time it takes to transmit a full frame. Whenever a node has a frame to send and it senses an empty (idle) slot, it transmits with probability p and defers until the next slot with probability q = 1− p. If that next slot is also empty, the node again decides to transmit or defer, with probabilities p and q, respectively. If that next slot is not empty—that is, some other station has decided to transmit—then the node simply waits for the next idle slot and the algorithm repeats. Returning to our discussion of the Ethernet, because there is no centralized con- trol it is possible for two (or more) adaptors to begin transmitting at the same time, either because both found the line to be idle or because both had been waiting for a busy line to become idle. When this happens, the two (or more) frames are said to collide on the network. Each sender, because the Ethernet supports collision detec- tion, is able to determine that a collision is in progress. At the moment an adaptor detects that its frame is colliding with another, it first makes sure to transmit a 32-bit jamming sequence and then stops the transmission. Thus, a transmitter will mini- mally send 96 bits in the case of a collision: 64-bit preamble plus 32-bit jamming sequence. One way that an adaptor will send only 96 bits—which is sometimes called a runt frame—is if the two hosts are close to each other. Had the two hosts been farther apart, they would have had to transmit longer, and thus send more bits, before detecting the collision. In fact, the worst-case scenario happens when the two hosts are at opposite ends of the Ethernet. To know for sure that the frame it just sent did not collide with another frame, the transmitter may need to send as many as 512 bits. Not coincidentally, every Ethernet frame must be at least 512 bits (64 bytes) long: 14 bytes of header plus 46 bytes of data plus 4 bytes of CRC. Why 512 bits? The answer is related to another question you might ask about an Ethernet: Why is its length limited to only 2500 m? Why not 10 or 1000 km? The answer to both questions has to do with the fact that the farther apart two nodes are, the longer it takes for a frame sent by one to reach the other, and the network is vulnerable to a collision during this time. 118 2 Direct Link Networks (a) (b) (c) AB AB AB A B (d) Figure 2.28 Worst-case scenario: (a) A sends a frame at time t; (b) A’s frame arrives at B at time t + d; (c) B begins transmitting at time t + d and collides with A’s frame; (d) B’s runt (32-bit) frame arrives at A at time t +2d. Figure 2.28 illustrates the worst-case scenario, where hosts A and B are at op- posite ends of the network. Suppose host A begins transmitting a frame at time t,as shown in (a). It takes it one link latency (let’s denote the latency as d) for the frame to reach host B. Thus, the first bit of A’s frame arrives at B at time t + d, as shown in (b). Suppose an instant before host A’s frame arrives (i.e., B still sees an idle line), host B begins to transmit its own frame. B’s frame will immediately collide with A’s frame, and this collision will be detected by host B (c). Host B will send the 32-bit jamming sequence, as described above. (B’s frame will be a runt.) Unfortunately, host A will not know that the collision occurred until B’s frame reaches it, which will happen one link latency later, at time t +2×d, as shown in (d). Host A must continue to transmit until this time in order to detect the collision. In other words, host A must transmit for 2×d to be sure that it detects all possible collisions. Considering that a maximally config- ured Ethernet is 2500 m long, and that there may be up to four repeaters between any two hosts, the round-trip delay has been determined to be 51.2 μs, which on a 10-Mbps Ethernet corresponds to 512 bits. The other way to look at this situa- tion is that we need to limit the Ethernet’s maximum latency to a fairly small value 2.6 Ethernet (802.3) 119 (e.g., 51.2 μs) for the access algorithm to work; hence, an Ethernet’s maximum length must be something on the order of 2500 m. Once an adaptor has detected a collision and stopped its transmission, it waits a certain amount of time and tries again. Each time it tries to transmit but fails, the adaptor doubles the amount of time it waits before trying again. This strategy of doubling the delay interval between each retransmission attempt is a general technique known as exponential backoff. More precisely, the adaptor first delays either 0 or 51.2 μs, selected at random. If this effort fails, it then waits 0, 51.2, 102.4, or 153.6 μs (selected randomly) before trying again; this is k × 51.2 for k = 0..3. After the third collision, it waits k× 51.2 for k = 0..23 − 1, again selected at random. In general, the algorithm randomly selects a k between 0 and 2n − 1 and waits k × 51.2 μs, where n is the number of collisions experienced so far. The adaptor gives up after a given number of tries and reports a transmit error to the host. Adaptors typically retry up to 16 times, although the backoff algorithm caps n in the above formula at 10. 2.6.3 Experience with Ethernet Because Ethernets have been around for so many years and are so popular, we have a great deal of experience in using them. One of the most important observations people have made about Ethernets is that they work best under lightly loaded conditions. This is because under heavy loads—typically, a utilization of over 30% is considered heavy on an Ethernet—too much of the network’s capacity is wasted by collisions. Fortunately, most Ethernets are used in a far more conservative way than the standard allows. For example, most Ethernets have fewer than 200 hosts connected to them, which is far fewer than the maximum of 1024. (See if you can discover a reason for this upper limit of around 200 hosts in Chapter 4.) Similarly, most Ethernets are far shorter than 2500 m, with a round-trip delay of closer to 5 μs than 51.2 μs. Another factor that makes Ethernets practical is that, even though Ethernet adaptors do not implement link-level flow control, the hosts typically provide an end-to-end flow-control mechanism. As a result, it is rare to find situations in which any one host is continuously pumping frames onto the network. Finally, it is worth saying a few words about why Ethernets have been so suc- cessful, so that we can understand the properties we should emulate with any LAN technology that tries to replace it. First, an Ethernet is extremely easy to administer and maintain: There are no switches that can fail, no routing or configuration tables that have to be kept up-to-date, and it is easy to add a new host to the network. It is hard to imagine a simpler network to administer. Second, it is inexpensive: Cable is cheap, and the only other cost is the network adaptor on each host. Any switch-based approach will involve an investment in some relatively expensive infrastructure (the switches), in addition to the incremental cost of each adaptor. As we will see in the 120 2 Direct Link Networks next chapter, the most successful LAN switching technology in use today is itself based on Ethernet. 2.7 Token Rings (802.5, FDDI) Alongside the Ethernet, token rings are the other significant class of shared-media network. There are more different types of token rings than there are types of Ethernets; this section will discuss the type that was for years the most prevalent, known as the IBM Token Ring. Like the Xerox Ethernet, IBM’s Token Ring has a nearly identical IEEE standard, known as 802.5. Where necessary, we note the differences between the IBM and 802.5 token rings. Most of the general principles of token ring networks can be understood once the IBM and 802.5 standards have been discussed. However, the FDDI (Fiber Dis- tributed Data Interface) standard—a newer, faster type of token ring—warrants some discussion, which we provide at the end of this section. At the time of writing, yet another token ring standard, called Resilient Packet Ring or 802.17, is nearing com- pletion. As the name suggests, a token ring network consists of a set of nodes connected in a ring (see Figure 2.29). Data always flows in a particular direction around the ring, with each node receiving frames from its upstream neighbor and then forwarding them to its downstream neighbor. This ring-based topology is in contrast to the Ethernet’s bus topology. Like the Ethernet, however, the ring is viewed as a single shared medium; it does not behave as a collection of independent point-to-point links that just happen Figure 2.29 Token ring network. 2.7 Token Rings (802.5, FDDI) 121 to be configured in a loop. Thus, a token ring shares two key features with an Ethernet: First, it involves a distributed algorithm that controls when each node is allowed to transmit, and second, all nodes see all frames, with the node identified in the frame header as the destination saving a copy of the frame as it flows past. The word “token” in token ring comes from the way access to the shared ring is managed. The idea is that a token, which is really just a special sequence of bits, circulates around the ring; each node receives and then forwards the token. When a node that has a frame to transmit sees the token, it takes the token off the ring (i.e., it does not forward the special bit pattern) and instead inserts its frame into the ring. Each node along the way simply forwards the frame, with the desti- nation node saving a copy and forwarding the message onto the next node on the ring. When the frame makes its way back around to the sender, this node strips its frame off the ring (rather than continuing to forward it) and reinserts the token. In this way, some node downstream will have the opportunity to transmit a frame. The media access algorithm is fair in the sense that as the token circulates around the ring, each node gets a chance to transmit. Nodes are serviced in a round-robin fashion. 2.7.1 Physical Properties One of the first things you might worry about with a ring topology is that any link or node failure would render the whole network useless. This problem is addressed by connecting each station into the ring using an electromechanical relay. As long as the station is healthy, the relay is open and the station is included in the ring. If the station stops providing power, the relay closes and the ring automatically bypasses the station. This is illustrated in Figure 2.30. Host From previous host To next host Relay (a) Host Host Host From previous host To next host Relay (b) Figure 2.30 Relay used on a token ring: (a) relay open—host active; (b) relay closed— host bypassed. 122 2 Direct Link Networks Host Host Host Host From previous MSAU To next MSAU MSAU Figure 2.31 Multistation access unit. Several of these relays are usually packed into a single box, known as a multi- station access unit (MSAU). This has the interesting effect of making a token ring actually look more like a star topology, as shown in Figure 2.31. It also makes it very easy to add stations to and remove stations from the network, since they can just be plugged into or unplugged from the nearest MSAU, while the overall wiring of the network can be left unchanged. One of the small differences between the IBM Token Ring specification and 802.5 is that the former actually requires the use of MSAUs, while the latter does not. In practice, MSAUs are almost always used because of the need for robustness and ease of station addition and removal. There are a few other physical details to know about 802.5 and IBM Token Rings. The data rate may be either 4 Mbps or 16 Mbps. The encoding of bits uses differential Manchester encoding, as described in Section 2.2. IBM Token Rings may have up to 260 stations per ring, while 802.5 sets the limit at 250. The physical medium is twisted pair for IBM, but is not specified in 802.5. 2.7.2 Token Ring Media Access Control It is now time to look a little more closely at how the MAC protocol operates on a token ring. The network adaptor for a token ring contains a receiver, a transmitter, and one or more bits of data storage between them. When none of the stations connected to the ring has anything to send, the token circulates around the ring. Obviously, the ring has to have enough “storage capacity” to hold an entire token. For example, the 2.7 Token Rings (802.5, FDDI) 123 802.5 token is 24 bits long. If every station could hold only 1 bit (as is the norm for 802.5 networks), and the stations were close enough together that the time for a bit to propagate from one station to another was negligible, we would need to have at least 24 stations on the ring before it would operate correctly. This situation is avoided by having one designated station, called the monitor, add some additional bits of delay to the ring if necessary. The operation of the monitor is described in more detail below. As the token circulates around the ring, any station that has data to send may “seize” the token, that is, drain it off the ring and begin sending data. In 802.5 net- works, the seizing process involves simply modifying 1 bit in the second byte token; the first 2 bytes of the modified token now become the preamble for the subsequent data packet. Once a station has the token, it is allowed to send one or more packets—exactly how many more depends on some factors described below. Each transmitted packet contains the destination address of the intended receiver; it may also contain a multicast (or broadcast) address if it is intended to reach more than one (or all) receivers. As the packet flows past each node on the ring, each node looks inside the packet to see if it is the intended recipient. If so, it copies the packet into a buffer as it flows through the network adaptor, but it does not remove the packet from the ring. The sending station has the responsibility of removing the packet from the ring. For any packet that is longer than the number of bits that can be stored in the ring, the sending station will be draining the first part of the packet from the ring while still transmitting the latter part. One issue we must address is how much data a given node is allowed to transmit each time it possesses the token, or said another way, how long a given node is allowed to hold the token. We call this the token holding time (THT). If we assume that most nodes on the network do not have data to send at any given time—a reasonable assumption, and certainly one that the Ethernet takes advantage of—then we could make a case for letting a node that possesses the token transmit as much data as it has before passing the token on to the next node. This would mean setting the THT to infinity. It would be silly in this case to limit a node to sending a single message and to force it to wait until the token circulates all the way around the ring before getting a chance to send another message. Of course, “as much data as it has” would be dangerous because a single station could keep the token for an arbitrarily long time, but we could certainly set the THT to significantly more than the time to send one packet. It is easy to see that the more bytes a node can send each time it has the token, the better the utilization of the ring you can achieve in the situation in which only a single node has data to send. The downside, of course, is that this strategy does not work well when multiple nodes have data to send—it favors nodes that have a lot of data to send over nodes that have only a small message to send, even when it is important 124 2 Direct Link Networks to get this small message delivered as soon as possible. The situation is analogous to finding yourself in line at the bank behind a customer who is taking out a car loan, even though you simply want to cash a check. In 802.5 networks, the default THT is 10 ms. There is a little subtlety to the use of the THT. Before putting each packet onto the ring, the station must check that the amount of time it would take to transmit the packet would not cause it to exceed the token holding time. This means keeping track of how long it has already held the token, and looking at the length of the next packet that it wants to send. From the token holding time we can derive another useful quantity, the token rotation time (TRT), which is the amount of time it takes a token to traverse the ring as viewed by a given node. It is easy to see that TRT ≤ ActiveNodes × THT + RingLatency where RingLatency denotes how long it takes the token to circulate around the ring when no one has data to send, and ActiveNodes denotes the number of nodes that have data to transmit. The 802.5 protocol provides a form of reliable delivery using 2 bits in the packet trailer, the A and C bits. These are both 0 initially. When a station sees a frame for which it is the intended recipient, it sets the A bit in the frame. When it copies the frame into its adaptor, it sets the C bit. If the sending station sees the frame come back over the ring with the A bit still 0, it knows that the intended recipient is not functioning or absent. If the A bit is set but not the C bit, this implies that for some reason (e.g., lack of buffer space) the destination could not accept the frame. Thus, the frame might reasonably be retransmitted later in the hope that buffer space had become available. Another detail of the 802.5 protocol concerns the support of different levels of priority. The token contains a 3-bit priority field, so we can think of the token having a certain priority n at any time. Each device that wants to send a packet assigns a priority to that packet, and the device can only seize the token to transmit a packet if the packet’s priority is at least as great as the token’s. The priority of the token changes over time due to the use of three reservation bits in the frame header. For example, a station X waiting to send a priority n packet may set these bits to n if it sees a data frame going past and the bits have not already been set to a higher value. This causes the station that currently holds the token to elevate its priority to n when it releases it. Station X is responsible for lowering the token priority to its old value when it is done. Note that this is a strict priority scheme, in the sense that no lower-priority pack- ets get sent when higher-priority packets are waiting. This may cause lower-priority packets to be locked out of the ring for extended periods if there is a sufficient supply of high-priority packets. 2.7 Token Rings (802.5, FDDI) 125 Token Frame Token Frame (a) (b) Figure 2.32 Token release: (a) early versus (b) delayed. One final issue will complete our discussion of the MAC protocol, which is the matter of exactly when the sending node releases the token. As illustrated in Figure 2.32, the sender can insert the token back onto the ring immediately following its frame (this is called early release) or after the frame it transmits has gone all the way around the ring and been removed (this is called delayed release). Clearly, early release allows better bandwidth utilization, especially on large rings. 802.5 originally used delayed token release, but support for early release was subsequently added. 2.7.3 Token Ring Maintenance As we noted above, token rings have a designated monitor station. The monitor’s job is to ensure the health of the ring. Any station on the ring can become the monitor, and there are defined procedures by which the monitor is elected when the ring is first connected or on the failure of the current monitor. A healthy monitor periodically announces its presence with a special control message; if a station fails to see such a message for some period of time, it will assume that the monitor has failed and will try to become the monitor. The procedures for electing a monitor are the same whether the ring has just come up or the active monitor has just failed. When a station decides that a new monitor is needed, it transmits a “claim token” frame, announcing its intent to become the new monitor. If that token circulates back to the sender, it can assume that it is OK for it to become the monitor. If some other station is also trying to become the monitor at the same instant, the sender might see a claim token message from that other station first. In this case, it will be necessary to break the tie using some well-defined rule like “highest address wins.” 126 2 Direct Link Networks Once the monitor is agreed upon, it plays a number of roles. We have already seen that it may need to insert additional delay into the ring. It is also responsible for making sure that there is always a token somewhere in the ring, either circulating or currently held by a station. It should be clear that a token may vanish for several reasons, such as a bit error, or a crash on the part of a station that was holding it. To detect a missing token, the monitor watches for a passing token and maintains a timer equal to the maximum possible token rotation time. This interval equals NumStations × THT + RingLatency where NumStations is the number of stations on the ring, and RingLatency is the total propagation delay of the ring. If the timer expires without the monitor seeing a token, it creates a new one. The monitor also checks for corrupted or orphaned frames. The former have checksum errors or invalid formats, and without monitor intervention, they could circulate forever on the ring. The monitor drains them off the ring before reinserting the token. An orphaned frame is one that was transmitted correctly onto the ring but whose “parent” died; that is, the sending station went down before it could remove the frame from the ring. These are detected using another header bit, the “‘monitor” bit. This is 0 on transmission and set to 1 the first time the packet passes the monitor. If the monitor sees a packet with this bit set, it knows the packet is going by for the second time and it drains the packet off the ring. One additional ring maintenance function is the detection of dead stations. The relays in the MSAU can automatically bypass a station that has been disconnected or powered down, but may not detect more subtle failures. If any station suspects a failure on the ring, it can send a beacon frame to the suspect destination. Based on how far this frame gets, the status of the ring can be established, and malfunctioning stations can be bypassed by the relays in the MSAU. 2.7.4 Frame Format We are now ready to define the 802.5 frame format, which is depicted in Figure 2.33. As noted above, 802.5 uses differential Manchester encoding. This fact is used by the frame format, which uses “illegal” Manchester codes in the start and end delimiters. Body ChecksumSrc addr Variable48 Dest addr 48 32 End delimiter 8 Frame status 8 Frame control 8 Access control 8 Start delimiter 8 Figure 2.33 802.5/token ring frame format. 2.7 Token Rings (802.5, FDDI) 127 After the start delimiter comes the access control byte, which includes the frame priority and the reservation priority mentioned above. The frame control byte is a demux key that identifies the higher-layer protocol. Similar to the Ethernet, 802.5 addresses are 48 bits long. The standard actually allows for smaller 16-bit addresses, but 48-bit addresses are typically used. When 48- bit addresses are used, they are interpreted in exactly the same way as on an Ethernet. The frame also includes a 32-bit CRC. This is followed by the frame status byte, which includes the A and C bits for reliable delivery. 2.7.5 FDDI In many respects, FDDI is similar to 802.5 and IBM Token Rings. However, there are significant differences—some arising because it runs on fiber, not copper, and some arising from innovations that were made subsequent to the invention of the IBM Token Ring. We discuss some of the significant differences below. Physical Properties Unlike 802.5 networks, an FDDI network consists of a dual ring—two independent rings that transmit data in opposite directions, as illustrated in Figure 2.34(a). The second ring is not used during normal operation but instead comes into play only if the primary ring fails, as depicted in Figure 2.34(b). That is, the ring loops back on the secondary fiber to form a complete ring, and as a consequence, an FDDI network is able to tolerate a single break in the cable or the failure of one station. Because of the expense of the dual-ring configuration, FDDI allows nodes to attach to the network by means of a single cable. Such nodes are called single (a) (b) Figure 2.34 Dual-fiber ring: (a) normal operation; (b) failure of the primary ring. 128 2 Direct Link Networks Concentrator (DAS) SAS SAS SAS SAS Upstream neighbor Downstream neighbor Figure 2.35 SASs connected to a concentrator. attachment stations (SAS); their dual-connected counterparts are called, not surpris- ingly, dual attachment stations (DAS). A concentrator is used to attach several SASs to the dual ring, as illustrated in Figure 2.35. Notice how the single-cable (two-fiber) connection into an SAS forms a connected piece of the ring. Should this SAS fail, the concentrator detects this situation and uses an optical bypass to isolate the failed SAS, thereby keeping the ring connected. This is analogous to the relays inside MSAUs used in 802.5 rings. Note that in this illustration, the second (backup) ring is denoted with a dotted line. As in 802.5, each network adaptor holds some number of bits between its input and output interfaces. Unlike 802.5, however, the buffer can be of different sizes in different stations, although never less than 9 bits nor more than 80 bits. It is also possible for a station to start transmitting bits out of this buffer before it is full. Of course, the total time it takes for a token to pass around the network is a function of the size of these buffers. For example, because FDDI is a 100-Mbps network, it has a 10-nanosecond (ns) bit time (each bit is 10 ns wide). If each station implements a 10-bit buffer and waits for the buffer to be half full before starting to transmit, then each station introduces a 5 × 10 ns = 50-ns delay into the total ring rotation time. FDDI has other physical characteristics. For example, the standard limits a single network to at most 500 stations (hosts), with a maximum distance of 2 km between any pair of stations. Overall, the network is limited to a total of 200 km of fiber, which means that, because of the dual nature of the ring, the total amount of cable connecting all stations is limited to 100 km. Also, although the “F” in FDDI implies that optical fiber serves as the underlying physical medium, the standard has been defined to run over a number of different physical media, including coax and twisted pair. Of course, you still have to be careful about the total distance covered by the ring. As we will 2.7 Token Rings (802.5, FDDI) 129 see below, the amount of time it takes the token to traverse the network plays an important role in the access control algorithm. FDDI uses 4B/5B encoding, as discussed in Section 2.2. Since FDDI was the first popular networking technology to use fiber, and 4B/5B chip sets operating at FDDI rates became widely available, 4B/5B has enjoyed considerable popularity as an encoding scheme for fiber. Timed Token Algorithm The rules governing token holding times are a little more complex in FDDI than in 802.5. The THT for each node is defined as before and is configured to some suitable value. In addition, to ensure that a given node has the opportunity to transmit within a certain amount of time—that is, to put an upper bound on the TRT observed by any node—we define a target token rotation time (TTRT), and all nodes agree to live within the limits of the TTRT. (How the nodes agree to a particular TTRT is described in the next subsection.) Specifically, each node measures the time between successive arrivals of the token. We call this the node’s measured TRT. If this measured TRT is greater than the agreed-upon TTRT, then the token is late, and the node does not transmit any data. If this measured TRT is less than the TTRT, then the token is early, and the node is allowed to hold the token for the difference between TTRT and the measured TRT. Although it may seem that we are now done, the algorithm we have just developed does not ensure that a node concerned with sending a frame with a bounded delay will actually be able to do so. The problem is that a node with lots of data to send has the opportunity, upon seeing an early token, to hold the token for so long that by the time a downstream node gets the token, its measured TRT is equal to or exceeds the TTRT, meaning that it still cannot transmit its frame. To account for this possibility, FDDI defines two classes of traffic: synchronous and asynchronous.3 When a node receives a token, it is always allowed to send synchronous data, without regard for whether the token is early or late. In contrast, a node can send asynchronous traffic only when the token is early. Note that the terms synchronous and asynchronous are somewhat misleading. By synchronous, FDDI means that the traffic is delay sensitive. For example, you would send voice or video as synchronous traffic on an FDDI network. In contrast, asynchronous means that the application is more interested in throughput than delay. A file transfer application would be asynchronous FDDI traffic. Are we done yet? Not quite. Because synchronous traffic can transmit without regard to whether the token is early or late, it would seem that if each node had a sizable 3Originally, FDDI defined two subclasses of asynchronous traffic: restricted and unrestricted. In practice, however, the restricted asynchronous case is not supported, and so we describe only the unrestricted case and refer to it simply as “asynchronous.” 130 2 Direct Link Networks amount of synchronous data to send, then the target rotation time would again be meaningless. To account for this, the total amount of synchronous data that can be sent during one token rotation is also bounded by TTRT. This means that in the worst case, the nodes with asynchronous traffic first use up one TTRT’s worth of time, and then the nodes with synchronous data consume another TTRT’s worth of time, meaning that it is possible for the measured TRT at any given node to be as much as 2 × TTRT. Note that if the synchronous traffic has already consumed one TTRT’s worth of time, then the nodes with asynchronous traffic will not send any data because the token will be late. Thus, while it is possible for a single rotation of the token to take as long as 2 × TTRT, it is not possible to have back-to-back rotations that take 2 × TTRT amount of time. One final detail concerns precisely how a node determines if it can send asyn- chronous traffic. As stated above, a node sends if the measured TRT is less than the TTRT. The question then arises: What if the measured TRT is less than the TTRT, but by such a small amount that it’s not possible to send the full message without exceeding the TTRT? The answer is that the node is allowed to send in this case. As a consequence, the measured TRT is actually bounded by TTRT plus the time it takes to send a full FDDI frame. Token Maintenance The FDDI mechanisms for ensuring that a valid token is always in circulation are also different from those in 802.5, as they are intertwined with the process of setting the TTRT. First, all nodes on an FDDI ring monitor the ring to be sure that the token has not been lost. Observe that in a correctly functioning ring, each node should see a valid transmission—either a data frame or the token—every so often. The greatest idle time between valid transmissions that a given node should experience is equal to the ring latency plus the time it takes to transmit a full frame, which on a maximally sized ring is a little less than 2.5 ms. Therefore, each node sets a timer event that fires after 2.5 ms. If this timer expires, the node suspects that something has gone wrong and transmits a “claim” frame. Every time a valid transmission is received, however, the node resets the timer back to 2.5 ms. The claim frames in FDDI differ from those in 802.5 because they contain the node’s bid for the TTRT, that is, the token rotation time that the node needs so that the applications running on the node can meet their timing constraints. A node can send a claim frame without holding the token and typically does so whenever it suspects a failure or when it first joins the network. If this claim frame makes it all the way around the ring, then the sender removes it, knowing that its TTRT bid was the lowest. That node now holds the token—that is, it is responsible for inserting a valid token on the ring—and may proceed with the normal token algorithm. 2.8 Wireless (802.11) 131 Control 88 824 CRCStart of frame End of frame Dest addr Body 4848 Src addr Status 32 Figure 2.36 FDDI frame format. When a node receives a claim frame, it checks to see if the TTRT bid in the frame is less than its own. If it is, then the node resets its local definition of the TTRT to that contained in the claim frame and forwards the frame to the next node. If the bid TTRT is greater than that node’s minimum required TTRT, then the claim frame is removed from the ring and the node enters the bidding process by putting its own claim frame on the ring. Should the bid TTRT be equal to the node’s required TTRT, the node compares the address of the claim frame’s sender with its own and the higher address wins. Thus, if a claim frame makes it all the way back around to the original sender, that node knows that it is the only active bidder and that it can safely claim the token. At the same time, all nodes are now in agreement about the TTRT that will be short enough to keep all nodes happy. Frame Format The FDDI frame format, depicted in Figure 2.36, differs in very few ways from that for 802.5. Because FDDI uses 4B/5B encoding instead of Manchester, it uses 4B/5B control symbols rather than illegal Manchester symbols in the start- and end-of-frame markers. The other significant differences are the presence of a bit in the header to distinguish synchronous from asynchronous traffic, and the lack of the access control bits of 802.5. 2.8 Wireless (802.11) Wireless networking is a rapidly evolving technology for connecting computers. As we saw earlier in this chapter, the possibilities for building wireless networks are almost endless, ranging from using infrared signals within a single building to constructing a global network from a grid of low-orbit satellites. This section takes a closer look at a specific technology centered around the emerging IEEE 802.11 standard. Like its Ethernet and token ring siblings, 802.11 is designed for use in a limited geographical area (homes, office buildings, campuses), and its primary challenge is to mediate access to a shared communication medium—in this case, signals propagating through space. 802.11 supports additional features (e.g., time-bounded services, power management, and security mechanisms), but we focus our discussion on its base functionality. 132 2 Direct Link Networks 2.8.1 Physical Properties 802.11 was designed to run over three different physical media—two based on spread spectrum radio and one based on diffused infrared. The radio-based versions currently run at 11 Mbps, but may soon run at 54 Mbps. The idea behind spread spectrum is to spread the signal over a wider frequency band than normal, so as to minimize the impact of interference from other devices. (Spread spectrum was originally designed for military use, so these “other devices” were often attempting to jam the signal.) For example, frequency hopping is a spread spectrum technique that involves transmitting the signal over a random sequence of frequencies; that is, first transmitting at one frequency, then a second, then a third, and so on. The sequence of frequencies is not truly random, but is instead computed algorithmically by a pseudorandom number generator. The receiver uses the same algorithm as the sender—and initializes it with the same seed—and hence is able to hop frequencies in sync with the transmitter to correctly receive the frame. A second spread spectrum technique, called direct sequence, achieves the same effect by representing each bit in the frame by multiple bits in the transmitted signal. For each bit the sender wants to transmit, it actually sends the exclusive-OR of that bit and n random bits. As with frequency hopping, the sequence of random bits is generated by a pseudorandom number generator known to both the sender and the receiver. The transmitted values, known as an n-bit chipping code, spread the signal across a frequency band that is n times wider than the frame would have otherwise required. Figure 2.37 gives an example of a 4-bit chipping sequence. 802.11 defines one physical layer using frequency hopping (over 79 1-MHz-wide frequency bandwidths) and a second using direct sequence (using an 11-bit chipping sequence). Both standards run in the 2.4-GHz frequency band of the electromagnetic spectrum. In both cases, spread spectrum also has the interesting characteristic of making the signal look like noise to any receiver that does not know the pseudorandom sequence. Random sequence: 0100101101011001 Data stream: 1010 XOR of the two: 1011101110101001 0 0 0 1 1 1 Figure 2.37 Example 4-bit chipping sequence. 2.8 Wireless (802.11) 133 ABC D Figure 2.38 Example wireless network. The third physical standard for 802.11 is based on infrared signals. The trans- mission is diffused, meaning that the sender and receiver do not have to be aimed at each other and do not need a clear line of sight. This technology has a range of up to about 10 m and is limited to the inside of buildings only. 2.8.2 Collision Avoidance At first glance, it might seem that a wireless protocol would follow exactly the same algorithm as the Ethernet—wait until the link becomes idle before transmitting and back off should a collision occur—and to a first approximation, this is exactly what 802.11 does. The problem is more complicated in a wireless network, however, because not all nodes are always within reach of each other. Consider the situation depicted in Figure 2.38, where each of four nodes is able to send and receive signals that reach just the nodes to its immediate left and right. For example, B can exchange frames with A and C but it cannot reach D, while C can reach B and D but not A. (A and D’s reach is not shown in the figure.) Suppose both A and C want to communicate with B and so they each send it a frame. A and C are unaware of each other since their signals do not carry that far. These two frames collide with each other at B, but unlike an Ethernet, neither A nor C is aware of this collision. A and C are said to be hidden nodes with respect to each other. A related problem, called the exposed node problem, occurs under the following circumstances. Suppose B is sending to A in Figure 2.38. Node C is aware of this communication because it hears B’s transmission. It would be a mistake for C to conclude that it cannot transmit to anyone just because it can hear B’s transmission. For example, suppose C wants to transmit to node D. This is not a problem since C’s transmission to D will not interfere with A’s ability to receive from B. (It would interfere with A sending to B, but B is transmitting in our example.) 134 2 Direct Link Networks 802.11 addresses these two problems with an algorithm called Multiple Access with Collision Avoidance (MACA). The idea is for the sender and receiver to exchange control frames with each other before the sender actually transmits any data. This exchange informs all nearby nodes that a transmission is about to begin. Specifically, the sender transmits a Request to Send (RTS) frame to the receiver; the RTS frame includes a field that indicates how long the sender wants to hold the medium (i.e., it specifies the length of the data frame to be transmitted). The receiver then replies with a Clear to Send (CTS) frame; this frame echoes this length field back to the sender. Any node that sees the CTS frame knows that it is close to the receiver, and therefore cannot transmit for the period of time it takes to send a frame of the specified length. Any node that sees the RTS frame but not the CTS frame is not close enough to the receiver to interfere with it, and so is free to transmit. There are two more details to complete the picture. First, the receiver sends an ACK to the sender after successfully receiving a frame. All nodes must wait for this ACK before trying to transmit.4 Second, should two or more nodes detect an idle link and try to transmit an RTS frame at the same time, their RTS frames will collide with each other. 802.11 does not support collision detection, but instead the senders realize the collision has happened when they do not receive the CTS frame after a period of time, in which case they each wait a random amount of time before trying again. The amount of time a given node delays is defined by the same exponential backoff algorithm used on the Ethernet (see Section 2.6.2). 2.8.3 Distribution System As described so far, 802.11 would be suitable for an ad hoc configuration of nodes that may or may not be able to communicate with all other nodes, depending on how far apart they are. Moreover, since one of the advantages of a wireless network is that nodes are free to move around—they are not tethered by wire—the set of directly reachable nodes may change over time. To help deal with this mobility and partial connectivity, 802.11 defines additional structure on a set of nodes. Nodes are free to directly communicate with each other as just described, but in practice, they operate within this structure. Instead of all nodes being created equal, some nodes are allowed to roam (e.g., your laptop) and some are connected to a wired network infrastructure. The latter are called access points (AP), and they are connected to each other by a so-called distribution system. Figure 2.39 illustrates a distribution system that connects three access points, each of which services the nodes in some region. Each of these regions 4This ACK was not part of the original MACA algorithm, but was instead proposed in an extended version called MACAW: MACA for Wireless LANs. 2.8 Wireless (802.11) 135 B H A F G D AP-2 AP-3AP-1 CE Distribution system Figure 2.39 Access points connected to a distribution network. is analogous to a cell in a cellular phone system, with the APs playing the same role as a base station. The details of the distribution system are not important to this discussion—it could be an Ethernet or a token ring, for example. The only important point is that the distribution network runs at layer 2 of the ISO architecture; that is, it does not depend on any higher-level protocols. Although two nodes can communicate directly with each other if they are within reach of each other, the idea behind this configuration is that each node associates itself with one access point. For node A to communicate with node E, for example, A first sends a frame to its access point (AP-1), which forwards the frame across the distribution system to AP-3, which finally transmits the frame to E. How AP-1 knew to forward the message to AP-3 is beyond the scope of 802.11; it may have used the bridging protocol described in the next chapter (Section 3.2). What 802.11 does specify is how nodes select their access points and, more interestingly, how this algorithm works in light of nodes moving from one cell to another. The technique for selecting an AP is called scanning and involves the following four steps: 1 The node sends a Probe frame. 2 All APs within reach reply with a Probe Response frame. 3 The node selects one of the access points and sends that AP an Association Request frame. 4 The AP replies with an Association Response frame. 136 2 Direct Link Networks B H A F G D AP-2 AP-3AP-1 EC C Distribution system Figure 2.40 Node mobility. A node engages this protocol whenever it joins the network, as well as when it becomes unhappy with its current AP. This might happen, for example, because the signal from its current AP has weakened due to the node moving away from it. Whenever a node acquires a new AP, the new AP notifies the old AP of the change (this happens in step 4) via the distribution system. Consider the situation shown in Figure 2.40, where node C moves from the cell serviced by AP-1 to the cell serviced by AP-2. As it moves, it sends Probe frames, which eventually result in Probe Response frames from AP-2. At some point, C prefers AP-2 over AP-1, and so it associates itself with that access point. The mechanism just described is called active scanning since the node is actively searching for an access point. APs also periodically send a Beacon frame that advertises the capabilities of the access point; these include the transmission rates supported by the AP. This is called passive scanning, and a node can change to this AP based on the Beacon frame simply by sending it an Association Request frame back to the access point. 2.8.4 Frame Format Most of the 802.11 frame format, which is depicted in Figure 2.41, is exactly what we would expect. The frame contains the source and destination node addresses, each of which are 48 bits long; up to 2312 bytes of data; and a 32-bit CRC. The Control field contains three subfields of interest (not shown): a 6-bit Type field that indicates whether the frame carries data, is an RTS or CTS frame, or is being used by the 2.9 Network Adaptors 137 Addr4 PayloadSeqCtrlAddr3Addr2Addr1 CRC 0–18,4964816 32484848 Duration 16 Control 16 Figure 2.41 802.11 frame format. scanning algorithm; and a pair of 1-bit fields—called ToDS and FromDS—that are described below. The peculiar thing about the 802.11 frame format is that it contains four, rather than two, addresses. How these addresses are interpreted depends on the settings of the ToDS and FromDS bits in the frame’s Control field. This is to account for the possibility that the frame had to be forwarded across the distribution system, which would mean that the original sender is not necessarily the same as the most recent transmitting node. Similar reasoning applies to the destination address. In the simplest case, when one node is sending directly to another, both the DS bits are 0, Addr1 identifies the target node, and Addr2 identifies the source node. In the most complex case, both DS bits are set to 1, indicating that the message went from a wireless node onto the dis- tribution system, and then from the distribution system to another wireless node. With both bits set, Addr1 identifies the ultimate destination, Addr2 identifies the immediate sender (the one that forwarded the frame from the distribution system to the ultimate destination), Addr3 identifies the intermediate destination (the one that ac- cepted the frame from a wireless node and forwarded it across the distribution system), and Addr4 identifies the original source. In terms of the example given in Figure 2.39, Addr1 corresponds to E, Addr2 identifies AP-3, Addr3 corresponds to AP-1, and Addr4 identifies A. 2.9 Network Adaptors Nearly all the networking functionality described in this chapter is implemented in the network adaptor: framing, error detection, and the media access protocol. The only exceptions are the point-to-point automatic repeat request (ARQ) schemes described in Section 2.5, which are typically implemented in the lowest-level protocol running on the host. We conclude this chapter by describing the design of a generic network adaptor and the device driver software that controls it. When reading this section, keep in mind that no two network adaptors are exactly alike; they vary in countless small details. Our focus, therefore, is on their general characteristics, although we do include some examples from an actual adaptor to make the discussion more tangible. 138 2 Direct Link Networks Host I/O bus Adaptor Network linkBus interface Link interface Figure 2.42 Block diagram of a typical network adaptor. 2.9.1 Components A network adaptor serves as an interface between the host and the network, and as a result, it can be thought of as having two main components: a bus interface that understands how to communicate with the host and a link interface that speaks the correct protocol on the network. There must also be a communication path between these two components, over which incoming and outgoing data is passed. A simple block diagram of a network adaptor is depicted in Figure 2.42. Network adaptors are always designed for a specific I/O bus, which often pre- cludes moving an adaptor from one vendor’s machine to another.5 Each bus, in effect, defines a protocol that is used by the host’s CPU to program the adaptor, by the adap- tor to interrupt the host’s CPU, and by the adaptor to read and write memory on the host. One of the main features of an I/O bus is the data transfer rate that it supports. For example, a typical bus might have a 32-bit-wide data path (i.e., it can transfer 32 bits of data in parallel) running at 33 MHz (i.e., the bus’s cycle time is 33 ns), giving it a peak transfer rate of approximatley 1 Gbps, which would be enough to support a (unidirectional) 622-Mbps STS-12 link. Of course, the peak rate tells us almost nothing about the average rate, which may be much lower. The link-half of the adaptor implements the link-level protocol. For fairly mature technologies like Ethernet, the link-half of the adaptor is implemented by a chip set that can be purchased on the commodity market. For newer link technologies, how- ever, the link-level protocol may be implemented in software on a general-purpose microprocessor or perhaps with some form of programmable hardware, such as a field-programmable gate array (FPGA). These approaches generally add to the cost of 5Fortunately, there are standards in bus design just as there are in networking, so some adaptors can be used on machines from several vendors. 2.9 Network Adaptors 139 the adaptor but make it more flexible—it is easier to modify software than hardware and easier to reprogram FPGAs than to redesign boards. Because the host’s bus and the network link are, in all probability, running at different speeds, there is a need to put a small amount of buffering between the two halves of the adaptor. Typically, a small FIFO (byte queue) is enough to hide the asynchrony between the bus and the link. 2.9.2 View from the Host Since we have spent most of this chapter discussing various protocols that are imple- mented by the link-half of the adaptor, we now turn our attention to the host’s view of the network adaptor. Control Status Register A network adaptor, like any other device, is ultimately programmed by software run- ning on the CPU. From the CPU’s perspective, the adaptor exports a control status register (CSR) that is readable and writable from the CPU. The CSR is typically located at some address in the memory, thereby making it possible for the CPU to read and write just like any other memory location. The CPU writes to the CSR to instruct it to transmit and/or receive a frame and reads from the CSR to learn the current state of the adaptor. The following is an example CSR from the Lance Ethernet device, which is manufactured by Advanced Micro Devices (AMD). The Lance device actually has four different control status registers; the following shows the bit masks used to interpret the 16-bit CSR0. To set a bit on the adaptor, the CPU does an inclusive-OR of CSRO and the mask corresponding to the bit it wants to set. To determine if a particular bit is set, the CPU compares the AND of the contents of CSR0 and the mask against 0. /* * Control and status bits for CSR0. * * Legend: * RO - Read Only * RC - Read/Clear (writing 1 clears, writing 0 has no effect) * RW - Read/Write * W1 - Write-1-only (writing 1 sets, writing 0 has no effect) * RW1 - Read/Write-1-only (writing 1 sets, writing 0 has no effect) */ #define LE_ERR 0x8000 /* RO BABL | CERR | MISS | MERR */ #define LE_BABL 0x4000 /* RC transmitted too many bits */ #define LE_CERR 0x2000 /* RC No Heartbeat */ 140 2 Direct Link Networks #define LE_MISS 0x1000 /* RC Missed an incoming packet */ #define LE_MERR 0x0800 /* RC Memory Error; no acknowledge */ #define LE_RINT 0x0400 /* RC Received packet Interrupt */ #define LE_TINT 0x0200 /* RC Transmitted packet Interrupt */ #define LE_IDON 0x0100 /* RC Initialization Done */ #define LE_INTR 0x0080 /* RO BABL|MISS|MERR|RINT|TINT|IDON */ #define LE_INEA 0x0040 /* RW Interrupt Enable */ #define LE_RXON 0x0020 /* RO Receiver On */ #define LE_TXON 0x0010 /* RO Transmitter On */ #define LE_TDMD 0x0008 /* W1 Transmit Demand (send it now) */ #define LE_STOP 0x0004 /* RW1 Stop */ #define LE_STRT 0x0002 /* RW1 Start */ #define LE_INIT 0x0001 /* RW1 Initialize */ This definition says, for example, that the host writesa1totheleast significant bit of CSR0 (0x0001) to initialize the Lance chip. Similarly, if the host seesa1inthe sixth significant bit (0x0020) and in the fifth significant bit (0x0010), then it knows that the Lance chip is enabled to receive and transmit frames, respectively. Interrupts The host CPU could sit in a tight loop reading the adaptor’s control status register until something interesting happens and then take the appropriate action. On the Lance chip, for example, it could continually watch fora1inthe11th significant bit (0x0400), which would indicate that a frame has just arrived. This is called polling, and although it is not an unreasonable design in certain situations (e.g., a network router that has nothing better to do than wait for the next frame), it is not typically done on end hosts that could better spend their time running application programs. Instead of polling, most hosts only pay attention to the network device when the adaptor interrupts the host. The device raises an interrupt when an event that requires host intervention occurs—for example, a frame has been successfully transmitted or received, or an error occurred when the device was attempting to transmit or receive a frame. The host’s architecture includes a mechanism that causes a particular proce- dure inside the operating system to be invoked when such an interrupt occurs. This procedure is known as an interrupt handler, and it inspects the CSR to determine the cause of the interrupt and then takes the appropriate action. While servicing an interrupt, the host typically disables additional interrupts. This keeps the device driver from having to service multiple interrupts at one time. Because interrupts are disabled, the device driver must finish its job quickly (it does not have the time to execute the entire protocol stack), and under no circumstances can it afford to block (that is, suspend execution while awaiting some event). For example, this might be accomplished by having the interrupt handler dispatch a process to take 2.9 Network Adaptors 141 care of the frame and then return. Thus, the handler makes sure that the frame will get processed without having to spend valuable time actually processing the frame itself. Direct Memory Access versus Programmed I/O One of the most important issues in network adaptor design is how the bytes of a frame are transferred between the adaptor and the host memory. There are two basic mechanisms: direct memory access (DMA) and programmed I/O (PIO). With DMA, the adaptor directly reads and writes the host’s memory without any CPU involvement; the host simply gives the adaptor a memory address and the adaptor reads to (writes from) it. With PIO, the CPU is directly responsible for moving data between the adaptor and the host memory: To send a frame, the CPU sits in a tight loop that first reads a word from host memory and then writes it to the adaptor; to receive a frame, the CPU reads words from the adaptor and writes them to memory. We now consider DMA and PIO in more detail. When using DMA, there is no need to buffer frames on the adaptor; the adaptor reads and writes host memory. (A few bytes of buffering are needed to stage data between the bus and the link, as described above, but complete frames are not buffered on the adaptor.) The CPU is therefore responsible for giving the adaptor a pair of buffer descriptor lists: one to transmit out of and one to receive into. A buffer descriptor list is an array of address/length pairs, as illustrated in Figure 2.43. When receiving frames, the adaptor uses as many buffers as it needs to hold the incoming frame. For example, the descriptor illustrated in Figure 2.43 would cause Buffer descriptor list Memory buffers 100 1400 1500 1500 1500 … Figure 2.43 Buffer descriptor list. 142 2 Direct Link Networks an Ethernet adaptor that was attempting to receive a 1450-byte frame to put the first 100 bytes in the first buffer and the next 1350 bytes in the second buffer. If a sec- ond 1500-byte frame arrived immediately after the first, it would be placed entirely in the third buffer. That is, separate frames are placed in separate buffers, although a single frame may be scattered across mul- tiple buffers. This latter feature is usually called scatter-read. In practice, scatter-read is used when the network’s maximum frame size is so large that it is wasteful to allocate all buffers big enough to contain the largest possible arriving frame. An OS-specific mes- sage data structure, similar to the ones de- scribed in Section 1.4.3, would then be used to link together all the buffers that make up a single frame. Scatter-read is typically not used on an Ethernet because preallocat- ing 1500-byte buffers does not excessively waste memory. Output works in a similar way. When the host has a frame to transmit, it puts a pointer to the buffer that contains the frame in the transmit descriptor list. Devices that support gather-write allow the frame to be fragmented across multiple physical buffers. In practice, gather-write is more widely used than scatter-read because outgoing frames are often constructed in a piecemeal fashion, with more than one protocol contributing a buffer. For example, by the time a mes- sage makes it down the protocol stack and is ready to be transmitted, it consists of a buffer that contains the aggregate header (the collection of headers attached by var- ious protocols that processed the message) and a separate buffer that contains the ap- plication’s data. Frames, Buffers, and Messages As this section has suggested, the network adaptor is the place where the network comes in physical con- tact with the host. It also happens to be the place where three differ- ent worlds intersect: the network, the host architecture, and the host operating system. It turns out that each of these has a different termi- nology for talking about the same thing. It is important to recognize when this is happening. From the network’s perspec- tive, the adaptor transmits frames from the host and receives frames into the host. Most of this chapter has been presented from the net- work perspective, so you should have a good understanding of what the term “frame” means. From the perspective of the host architecture, each frame is re- ceived into or transmitted from a buffer, which is simply a region of main memory of some length and starting at some address. Fi- nally, from the operating system’s perspective, a message is an ab- stract object that holds network frames. Messages are implemented by a data structure that includes pointers to different memory loca- tions (buffers). We saw an exam- ple of a message data structure in Chapter 1. 2.9 Network Adaptors 143 Host Adaptor Memory CPU Memory Memory Figure 2.44 Programmed I/O. In the case of PIO, the network adaptor must contain some amount of buffering— the CPU copies frames between host memory and this adaptor memory, as illustrated in Figure 2.44. The basic fact that necessitates buffering is that, with most operating systems, you can never be sure when the CPU will get around to doing something, so you need to be prepared to wait for it. One important question that must be ad- dressed is how much memory is needed on the adaptor. There certainly needs to be at least one frame’s worth of memory in both the transmit and the receive direction. In addition, adaptors that use PIO usually have additional memory that can be used to hold a small number of incoming frames until the CPU can get around to copy- ing them into host memory. Although the computer system axiom that “memory is cheap” would seem to suggest putting a huge amount of memory on the adaptor, this memory must be of the more expensive dual-ported type because both the CPU and the adaptor read/write it. PIO-based adaptors typically have something on the order of 64–256 KB of adaptor memory, although there are adaptors with as much as 1 MB of memory. Device Drivers A device driver is a collection of operating system routines that effectively anchor the protocol stack to the network hardware. It typically includes routines to initialize the device, transmit frames on the link, and field interrupts. The code is often difficult to read because it’s full of device-specific details, but the overall logic is actually quite simple. For example, a transmit routine first makes sure there is a free transmit buffer on the device to handle the message. If not, it has to block the process until one is available. Once there is an available transmit buffer, the invoking process disables 144 2 Direct Link Networks interrupts to protect itself from interference. It then translates the message from the internal OS format to that expected by the device, sets the CSR to instruct the device to transmit, and enables interrupts. The logic for the interrupt handler is equally simple. It first disables additional interrupts that might interfere with the processing of this interrupt. It then inspects the CSR to determine what caused the interrupt. There are three possibilities: (1) an error has occurred, (2) a transmit request has completed, or (3) a frame has been received. In the first case, the handler prints a message and clears the error bits. In the second case, we know that a transmit request that was queued earlier by the transmit routine has completed, meaning that there is now a free transmit buffer that can be reused. In the third case, the handler calls a receive routine to extract the incoming frame from the receive buffer list and place it in the OS’s internal message data structure, and then start a process to shepherd the message up the protocol stack. 2.9.3 Memory Bottleneck As discussed in Section 2.1.1, host memory performance is often the limiting factor in network performance. Nowhere is this possibility more critical than at the host/adaptor interface. To help drive this point home, consider Figure 2.45. This diagram shows the bandwidth available between various components of a modern PC. While the I/O bus is fast enough to transfer frames between the network adaptor and host memory at gigabit rates, there are two potential problems. The first is that the advertised I/O bus speed corresponds to its peak bandwidth; it is the product of the bus’s width and clock speed (e.g., a 32-bit-wide bus running at 33 MHz has a peak transfer rate of 1056 Mbps). The real limitation is the size of the data block that is being transferred across the I/O bus, since there is a certain amount of overhead involved in each bus transfer. On some architectures, for example, I/O bus 235 Mbps 1056 Mbps CPU Memory Crossbar Figure 2.45 Memory bandwidth on a modern PC-class machine. 2.9 Network Adaptors 145 it takes 8 clock cycles to acquire the bus for the purpose of transferring data from the adaptor to host memory. This overhead is independent of the number of data bytes transferred. Thus, if you want to transfer a 64-byte payload across the I/O bus—this happens to be the size of a minimum Ethernet packet—then the whole transfer takes 24 cycles: 8 cycles to acquire the bus and 16 cycles to transfer the data. (The bus is 32 bits wide, which means that it can transfer a 4-byte word during each clock cycle; 64 bytes divided by 4 bytes per cycle equals 16 cycles.) This means that the maximum bandwidth you can achieve is 16 ÷ (8 + 16) × 1056 = 704 Mbps not the peak 1056 Mbps. The second problem is that the memory/CPU bandwidth, which is 235 MBps (1880 Mbps), is the same order of magnitude as the bandwidth of the I/O bus. Fortu- nately, this is a measured number rather than an advertised peak rate. The ramification is that while it is possible to deliver frames across the I/O bus and into memory and then to load the data from memory into the CPU’s registers at network bandwidths, it is impractical for the device driver, operating system, and application to go to mem- ory multiple times for each word of data in a network packet, possibly because it needs to copy the data from one buffer to another. In particular, if the memory/CPU path is crossed n times, then it might be the case that the bandwidth your applica- tion sees is 235 MBps/n. (The performance might be better if the data is cached, but often caches don’t help with data arriving from the network.) For example, if the various software layers need to copy the message from one buffer to another four times—not an uncommon situation—then the application might see a throughput of 58.75 MBps (470 Mbps), a far cry from the 1056 Mbps we thought this machine could support.◮ As an aside, it is important to recognize that there are many parallels between moving a message to and from memory and moving a message across a network. In particular, the effective throughput of the memory system is defined by the same two formulas given in Section 1.5. Throughput = TransferSize/TransferTime TransferTime = RTT + 1/Bandwidth × TransferSize In the case of the memory system, however, the transfer size corresponds to how big a unit of data we can move across the bus in one transfer (i.e., cache line versus small cells versus large message), and the RTT corresponds to the memory latency, that 146 2 Direct Link Networks is, whether the memory is on-chip cache, off-chip cache, or main memory. Just as in the case of the network, the larger the transfer size and the smaller the latency, the better the effective throughput. Also similar to a network, the effective memory throughput does not necessarily equal the peak memory bandwidth (i.e., the bandwidth that can be achieved with an infinitely large transfer). The main point of this discussion is that we must be aware of the limits memory bandwidth places on network performance. If carefully designed, the system can work around these limits. For example, it is possible to integrate the buffers used by the device driver, the operating system, and the application in a way that minimizes data copies. The system also needs to be aware of when data is brought into the cache, so it can perform all necessary operations on the data before it gets bumped from the cache. The details of how this is accomplished are beyond the scope of this book, but can be found in papers referenced at the end of the chapter. Finally, there is a second important lesson lurking in this discussion: when the network isn’t performing as well as you think it should, it’s not always the network’s fault. In many cases, the actual bottleneck in the system is one of the machines con- nected to the network. For example, when it takes a long time for a Web page to appear on your browser, it might be network congestion, but it’s just as likely the case that the server at the other end of the network can’t keep up with the workload. 2.10 Summary This chapter introduced the hardware building blocks of a computer network—nodes and links—and discussed the five key problems that must be solved so that two or more nodes that are directly connected by a physical link can exchange messages with each other. First, physical links carry signals. It is therefore necessary to encode the bits that make up a binary message into the signal at the source node and then to recover the bits from the signal at the receiving node. This is the encoding problem, and it is made challenging by the need to keep the sender’s and receiver’s clocks synchronized. We discussed four different encoding techniques—NRZ, NRZI, Manchester, and 4B/5B— which differ largely in how they encode clock information along with the data being transmitted. One of the key attributes of an encoding scheme is its efficiency, that is, the ratio of signal pulses to encoded bits. Once it is possible to transmit bits between nodes, the next step is to figure out how to package these bits into frames. This is the framing problem, and it boils down to being able to recognize the beginning and end of each frame. Again, we looked at several different techniques, including byte-oriented protocols, bit-oriented protocols, and clock-based protocols. Open Issue: Does It Belong in Hardware? 147 Assuming that each node is able to recognize the collection of bits that make up a frame, the third problem is to determine if those bits are in fact correct, or if they have possibly been corrupted in transit. This is the error detection problem, and we looked at three different approaches: cyclic redundancy check, two-dimensional parity, and checksums. Of these, the CRC approach gives the strongest guarantees and is the most widely used at the link level. Given that some frames will arrive at the destination node containing errors and thus will have to be discarded, the next problem is how to recover from such losses. The goal is to make the link appear reliable. The general approach to this problem is called ARQ and involves using a combination of acknowledgments and timeouts. We looked at three specific ARQ algorithms: stop-and-wait, sliding window, and concurrent channels. What makes these algorithms interesting is how effectively they use the link, with the goal being to keep the pipe full. The final problem is not relevant to point-to-point links, but it is the central issue in multiple-access links: how to mediate access to a shared link so that all nodes eventually have a chance to transmit their data. In this case, we looked at three dif- ferent media access protocols—Ethernet, token ring, and wireless—which have been put to practical use in building local area networks. What these technologies have in common is that control over the network is distributed over all the nodes connected to the network; there is no dependence on a central arbitrator. We concluded the chapter by observing that, in practice, most of the algorithms that address these five problems are implemented on the adaptor that connects the host to the link. It turns out that the design of this adaptor is of critical importance in how well the network, as a whole, performs. One of the most important questions in the design of any computer system is, What belongs in hardware and what belongs in software? In the case of networking, the network adaptor finds itself at the heart of this ques- tion. For example, why is the Ether- OPEN ISSUE Does It Belong in Hardware? net algorithm, presented in Section 2.6 of this chapter, typically implemented on the network adaptor, while the higher-level protocols discussed later in this book are not? It is certainly possible to put a general-purpose microprocessor on the network adaptor, which gives you the opportunity to move high-level protocols there, such as TCP/IP. The reason that this is typically not done is complicated, but it comes down to 148 2 Direct Link Networks the economics of computer design: The host processor is usually the fastest processor on a computer, and it would be a shame if this fast host processor had to wait for a slower adaptor processor to run TCP/IP when it could have done the job faster itself. On the flip side, some protocol processing does belong on the network adaptor. The general rule of thumb is that any processing for which a fixed processor can keep pace with the link speed—that is, a faster processor would not improve the situation—is a good candidate for being moved to the adaptor. In other words, any function that is already limited by the link speed, as opposed to the processor at the end of the link, might be effectively implemented on the adaptor. Historically, the decision as to what functionality belonged on the network adap- tor and what belonged on the host computer was a complex one that generated quite a body of research. In modern systems, it is almost always the case that the MAC layer and below are performed by the adaptor, while the IP layer and above are performed on the host. Interestingly, however, the same debate about how much hardware assis- tance is required above the MAC layer continues in the design of switches and routers, which is a topic for the next two chapters. Independent of exactly what protocols are implemented on the network adaptor, generally the data will eventually find its way onto the main computer, and when it does, the efficiency with which the data is moved between the adaptor and the com- puter’s memory is very important. Recall from Section 2.9.3 that memory bandwidth— the rate at which data can be moved from one memory location to another—has the potential to be a limiting factor in how a workstation-class machine performs. An inefficient host/adaptor data transfer mechanism can, therefore, limit the through- put rate seen by application programs running on the host. First, there is the issue of whether DMA or programmed I/O is used; each has advantages in different situations. Second, there is the issue of how well the network adaptor is integrated with the oper- ating system’s buffer mechanism; a carefully integrated system is usually able to avoid copying data at a higher level of the protocol graph, thereby improving application- to-application throughput. FURTHER READING One of the most important contributions in computer networking over the last 20 years is the original paper by Metcalf and Boggs (1976) introducing the Ethernet. Many years later, Boggs, Mogul, and Kent (1988) reported their practical experiences with Ethernet, debunking many of the myths that had found their way into the literature over the years. Both papers are must reading. The third and fourth papers discuss the issues involved in integrating high-speed network adaptors with system software. Further Reading 149 ■ Metcalf, R., and D. Boggs. Ethernet: Distributed packet switching for local computer networks. Communications of the ACM 19(7):395–403, July 1976. ■ Boggs, D., J. Mogul, and C. Kent. Measured capacity of an Ethernet. Pro- ceedings of the SIGCOMM ’88 Symposium, pages 222–234, August 1988. ■ Metcalf, R. Computer/network interface design lessons from Arpanet and Eth- ernet. IEEE Journal of Selected Areas in Communication (JSAC) 11(2):173– 180, February 1993. ■ Druschel, P., M. Abbot, M. Pagels, and L. L. Peterson. Network subsystem design. IEEE Network (Special Issue on End-System Support for High Speed Networks) 7(4):8–17, July 1993. There are countless textbooks with a heavy emphasis on the lower levels of the network hierarchy, with a particular focus on telecommunications—networking from the phone company’s perspective. Books by Spragins et al. [SHP91] and Minoli [Min93] are two good examples. Several other books concentrate on various local area network technologies. Of these, Stallings’s book is the most comprehensive [Sta00b], while Jain gives a thorough description of FDDI [Jai94]. Jain’s book also gives a good introduction to the low-level details of optical communication. Also, a comprehensive overview of FDDI can be found in Ross’s article [Ros86]. For an introduction to information theory, Blahut’s book is a good place to start [Bla87], along with Shannon’s seminal paper on link capacity [Sha48]. For a general introduction to the mathematics behind error codes, Rao and Fujiwara [RF89] is recommended. For a detailed discussion of the mathematics of CRCs in particular, along with some more information about the hardware used to calculate them, see Peterson and Brown [PB61]. On the topic of network adaptor design, much work was done in the early 1990s by researchers trying to connect hosts to networks running at higher and higher rates. In addition to the two examples given in the reading list, see Traw and Smith [TS93], Ramakrishnan [Ram93], Edwards et al. [EWL+94], Druschel et al. [DPD94], Kanakia and Cheriton [KC88], Cohen et al. [CFFD93], and Steenkiste [Ste94a]. Recently, a new generation of interface cards, ones that utilize network processors, are coming onto the market. Spalink et al. demonstrate how these new processors can be programmed to implement various network functionality [SKPG01]. For general information on computer architecture, Hennessy and Patterson’s book [HP02] is an excellent reference. Finally, we recommend the following live reference: ■ http://standards.ieee.org/: status of various IEEE network-related standards 150 2 Direct Link Networks EXERCISES 1 Show the NRZ, Manchester, and NRZI encodings for the bit pattern shown in Figure 2.46. Assume that the NRZI signal starts out low. 2 Show the 4B/5B encoding, and the resulting NRZI signal, for the following bit sequence: 1110 0101 0000 0011 3 Show the 4B/5B encoding, and the resulting NRZI signal, for the following bit sequence: 1101 1110 1010 1101 1011 1110 1110 1111 4 In the 4B/5B encoding (Table 2.4), only two of the 5-bit codes used end in two 0s. How many possible 5-bit sequences are there (used by the existing code or not) that meet the stronger restriction of having at most one leading and at most one trailing 0? Could all 4-bit sequences be mapped to such 5-bit sequences? 5 Assuming a framing protocol that uses bit stuffing, show the bit sequence trans- mitted over the link when the frame contains the following bit sequence: 110101111101011111101011111110 Mark the stuffed bits. 6 Suppose the following sequence of bits arrives over a link: 1101011111010111110010111110110 Bits NRZ 1001111100010001 Clock Manchester NRZI Figure 2.46 Diagram for Exercise 1. Exercises 151 Show the resulting frame after any stuffed bits have been removed. Indicate any errors that might have been introduced into the frame. 7 Suppose the following sequence of bits arrive over a link: 011010111110101001111111011001111110 Show the resulting frame after any stuffed bits have been removed. Indicate any errors that might have been introduced into the frame. 8 Suppose you want to send some data using the BISYNC framing protocol, and the last 2 bytes of your data are DLE and ETX. What sequence of bytes would be transmitted immediately prior to the CRC? 9 For each of the following framing protocols, give an example of a byte/bit sequence that should never appear in a transmission. (a) BISYNC (b) HDLC 10 Assume that a SONET receiver resynchronizes its clock whenever a 1 bit appears; otherwise, the receiver samples the signal in the middle of what it believes is the bit’s time slot. (a) What relative accuracy of the sender’s and receiver’s clocks is required in order to receive correctly 48 0 bytes (one ATM AAL5 cell’s worth) in a row? (b) Consider a forwarding station A on a SONET STS-1 line, receiving frames from the downstream end B and retransmitting them upstream. What relative accuracy of A’s and B’s clocks is required to keep A from accumulating more than one extra frame per minute? 11 Show that two-dimensional parity allows detection of all 3-bit errors. 12 Give an example of a 4-bit error that would not be detected by two-dimensional parity, as illustrated in Figure 2.16. What is the general set of circumstances under which 4-bit errors will be undetected? 13 Show that two-dimensional parity provides the receiver enough information to correct any 1-bit error (assuming the receiver knows only 1 bit is bad), but not any 2-bit error. 152 2 Direct Link Networks 14 Show that the Internet checksum will never be 0xFFFF (that is, the final value of sum will not be 0x0000) unless every byte in the buffer is 0. (Internet specifications in fact require that a checksum of 0x0000 be transmitted as 0xFFFF; the value 0x0000 is then reserved for an omitted checksum. Note that, in ones complement arithmetic, 0x0000 and 0xFFFF are both representations of the number 0.) 15 Prove the Internet checksum computation shown in the text is independent of byte order (host order or network order) except that the bytes in the final check- sum should be swapped later to be in the correct order. Specifically, show that the sum of 16-bit word integers can be computed in either byte order. For exam- ple, if the ones complement sum (denoted by +′) of 16-bit words is represented as [A, B] +′ [C, D] +′ ···+′ [Y, Z] the following swapped sum is the same as the original sum above: [B, A] +′ [D, C] +′ ···+′ [Z, Y] 16 Suppose that one byte in a buffer covered by the Internet checksum algorithm needs to be decremented (e.g., a header hop count field). Give an algorithm to compute the revised checksum without rescanning the entire buffer. Your algorithm should consider whether the byte in question is low order or high order. 17 Show that the Internet checksum can be computed by first taking the 32-bit ones complement sum of the buffer in 32-bit units, then taking the 16-bit ones com- plement sum of the upper and lower halfwords, and finishing as before by com- plementing the result. (To take a 32-bit ones complement sum on 32-bit twos complement hardware, you need access to the “overflow” bit.) 18 Suppose we want to transmit the message 11001001 and protect it from errors using the CRC polynomial x3 + 1. (a) Use polynomial long division to determine the message that should be trans- mitted. (b) Suppose the leftmost bit of the message is inverted due to noise on the trans- mission link. What is the result of the receiver’s CRC calculation? How does the receiver know that an error has occurred? Exercises 153 19 Suppose we want to transmit the message 1011 0010 0100 1011 and protect it from errors using the CRC-8 polynomial x8 + x2 + x1 + 1. (a) Use polynomial long division to determine the message that should be trans- mitted. (b) Suppose the leftmost bit of the message is inverted due to noise on the trans- mission link. What is the result of the receiver’s CRC calculation? How does the receiver know that an error has occurred? 20 The CRC algorithm as presented in this chapter requires lots of bit manipulations. It is, however, possible to do polynomial long division taking multiple bits at a time, via a table-driven method, that enables efficient software implementations of CRC. We outline the strategy here for long division 3 bits at a time (see Table 2.6); in practice we would divide 8 bits at a time, and the table would have 256 entries. Let the divisor polynomial C = C(x)bex3 + x2 + 1, or 1101. To build the table for C, we take each 3-bit sequence, p, append three trailing 0s, and then find the quotient q = p⌢000÷C, ignoring the remainder. The third column is the product C × q, the first 3 bits of which should equal p. (a) Verify, for p = 110, that the quotients p⌢000 ÷ C and p⌢111 ÷ C are the same; that is, it doesn’t matter what the trailing bits are. (b) Fill in the missing entries in the table. p q=p⌢000÷C C × q 000 000 000 000 001 001 001 101 010 011 010 011 0 011 100 111 100 011 101 110 101 110 110 100 110 111 111 Table 2.6 Table-driven CRC calculation. 154 2 Direct Link Networks (c) Use the table to divide 101 001 011 001 100 by C. Hint: The first 3 bits of the dividend are p = 101, so from the table the corresponding first 3 bits of the quotient are q = 110. Write the 110 above the second 3 bits of the divi- dend, and subtract C×q = 101 110, again from the table, from the first 6 bits of the dividend. Keep going in groups of 3 bits. There should be no remainder. 21 With 1 parity bit we can detect all 1-bit errors. Show that at least one generalization fails, as follows: (a) Show that if messages m are 8 bits long, then there is no error detection code e = e(m) of size 2 bits that can detect all 2-bit errors. Hint: Consider the set M of all 8-bit messages with a single 1 bit; note that any message from M can be transmuted into any other with a 2-bit error, and show that some pair of messages m1 and m2 in M must have the same error code e. (b) Find an N (not necessarily minimal) such that no 32-bit error detection code applied to N-bit blocks can detect all errors altering up to 8 bits. 22 Consider an ARQ protocol that uses only negative acknowledgments (NAKs), but no positive acknowledgments (ACKs). Describe what timeouts would need to be scheduled. Explain why an ACK-based protocol is usually preferred to a NAK-based protocol. 23 Consider an ARQ algorithm running over a 20-km point-to-point fiber link. (a) Compute the propagation delay for this link, assuming that the speed of light is 2 × 108 m/s in the fiber. (b) Suggest a suitable timeout value for the ARQ algorithm to use. (c) Why might it still be possible for the ARQ algorithm to time out and retransmit a frame, given this timeout value? 24 Suppose you are designing a sliding window protocol for a 1-Mbps point-to-point link to the moon, which has a one-way latency of 1.25 seconds. Assuming that each frame carries 1 KB of data, what is the minimum number of bits you need for the sequence number? 25 Suppose you are designing a sliding window protocol for a 1-Mbps point-to-point link to a stationary satellite revolving around the earth at 3 × 104 km altitude. Assuming that each frame carries 1 KB of data, what is the minimum number of bits you need for the sequence number in the following cases? Assume the speed of light is 3 × 108 m/s. Exercises 155 (a) RWS=1 (b) RWS = SWS 26 The text suggests that the sliding window protocol can be used to implement flow control. We can imagine doing this by having the receiver delay ACKs, that is, not send the ACK until there is free buffer space to hold the next frame. In doing so, each ACK would simultaneously acknowledge the receipt of the last frame and tell the source that there is now free buffer space available to hold the next frame. Explain why implementing flow control in this way is not a good idea. 27 Implicit in the stop-and-wait scenarios of Figure 2.19 is the notion that the re- ceiver will retransmit its ACK immediately on receipt of the duplicate data frame. Suppose instead that the receiver keeps its own timer and retransmits its ACK only after the next expected frame has not arrived within the timeout interval. Draw timelines illustrating the scenarios in Figure 2.19(b)–(d); assume the re- ceiver’s timeout value is twice the sender’s. Also redraw (c) assuming the receiver’s timeout value is half the sender’s. 28 In stop-and-wait transmission, suppose that both sender and receiver retransmit their last frame immediately on receipt of a duplicate ACK or data frame; such a strategy is superficially reasonable because receipt of such a duplicate is most likely to mean the other side has experienced a timeout. (a) Draw a timeline showing what will happen if the first data frame is somehow duplicated, but no frame is lost. How long will the duplications continue? This situation is known as the Sorcerer’s Apprentice bug. (b) Suppose that, like data, ACKs are retransmitted if there is no response within the timeout period. Suppose also that both sides use the same timeout interval. Identify a reasonably likely scenario for triggering the Sorcerer’s Apprentice bug. 29 Give some details of how you might augment the sliding window protocol with flow control by having ACKs carry additional information that reduces the SWS as the receiver runs out of buffer space. Illustrate your protocol with a timeline for a transmission; assume the initial SWS and RWS are 4, the link speed is instan- taneous, and the receiver can free buffers at the rate of one per second (i.e., the receiver is the bottleneck). Show what happens at T = 0, T = 1, ..., T = 4 sec- onds. 156 2 Direct Link Networks 30 Describe a protocol combining the sliding window algorithm with selective ACKs. Your protocol should retransmit promptly, but not if a frame simply arrives one or two positions out of order. Your protocol should also make explicit what happens if several consecutive frames are lost. 31 Draw a timeline diagram for the sliding window algorithm with SWS = RWS =3 frames for the following two situations. Use a timeout interval of about 2×RTT. (a) Frame 4 is lost. (b) Frames 4–6 are lost. 32 Draw a timeline diagram for the sliding window algorithm with SWS = RWS = 4 frames for the following two situations. Assume the receiver sends a duplicate acknowledgement if it does not receive the expected frame. For example, it sends DUPACK[2] when it expects to see FRAME[2] but receives FRAME[3] instead. Also, the receiver sends a cumulative acknowledgment after it receives all the outstand- ing frames. For example, it sends ACK[5] when it receives the lost frame FRAME[2] after it already received FRAME[3], FRAME[4], and FRAME[5]. Use a timeout inter- val of about 2 × RTT. (a) Frame 2 is lost. Retransmission takes place upon timeout (as usual). (b) Frame 2 is lost. Retransmission takes place either upon receipt of the first DUPACK or upon timeout. Does this scheme reduce the transaction time? Note that some end-to-end protocols (e.g., variants of TCP) use a similar scheme for fast retransmission. 33 Suppose that we attempt to run the sliding window algorithm with SWS = RWS =3 and with MaxSeqNum = 5. The Nth packet DATA[N] thus actually contains N mod 5 in its sequence number field. Give an example in which the algorithm becomes confused; that is, a scenario in which the receiver expects DATA[5] and accepts DATA[0]—which has the same transmitted sequence number—in its stead. No packets may arrive out of order.Note this implies MaxSeqNum ≥ 6 is necessary as well as sufficient. 34 Consider the sliding window algorithm with SWS = RWS = 3, with no out-of- order arrivals, and with infinite-precision sequence numbers. (a) Show that if DATA[6] is in the receive window, then DATA[0] (or in general any older data) cannot arrive at the receiver (and hence that MaxSeqNum = 6 would have sufficed). Exercises 157 ARB Figure 2.47 Diagram for Exercises 36–38. (b) Show that if ACK[6] may be sent (or, more literally, that DATA[5] is in the sending window), then ACK[2] (or earlier) cannot be received. These amount to a proof of the formula given in Section 2.5.2, particularized to the case SWS = 3. Note that part (b) implies that the scenario of the previous problem cannot be reversed to involve a failure to distinguish ACK[0] and ACK[5]. 35 Suppose that we run the sliding window algorithm with SWS = 5 and RWS = 3, and no out-of-order arrivals. (a) Find the smallest value for MaxSeqNum. You may assume that it suffices to find the smallest MaxSeqNum such that if DATA[MaxSeqNum] is in the receive window, then DATA[0] can no longer arrive. (b) Give an example showing that MaxSeqNum − 1 is not sufficient. (c) State a general rule for the minimum MaxSeqNum in terms of SWS and RWS. 36 Suppose A is connected to B via an intermediate router R, as shown in Figure 2.47. The A–R and R–B links each accept and transmit only one packet per second in each direction (so two packets take 2 seconds), and the two directions trans- mit independently. Assume A sends to B using the sliding window protocol with SWS = 4. (a) For Time = 0, 1, 2, 3, 4, 5, state what packets arrive at and leave each node, or label them on a timeline. (b) What happens if the links have a propagation delay of 1.0 seconds, but accept immediately as many packets as are offered (i.e., latency = 1 second but bandwidth is infinite)? 37 Suppose A is connected to B via an intermediate router R, as in the previous problem. The A–R link is instantaneous, but the R–B link transmits only one packet each second, one at a time (so two packets take 2 seconds). Assume A sends to B using the sliding window protocol with SWS = 4. For Time = 0, 1, 2, 3, 4, state what packets arrive at and are sent from A and B. How large does the queue at R grow? 158 2 Direct Link Networks 38 Consider the situation in the previous exercise, except this time assume that the router has a queue size of 1; that is, it can hold one packet in addition to the one it is sending (in each direction). Let A’s timeout be 5 seconds, and let SWS again be 4. Show what happens at each second from T = 0 until all four packets from the first windowful are successfully delivered. 39 Why is it important for protocols configured on top of the Ethernet to have a length field in their header, indicating how long the message is? 40 What kinds of problems can arise when two hosts on the same Ethernet share the same hardware address? Describe what happens and why that behavior is a problem. 41 The 1982 Ethernet specification allowed between any two stations up to 1500 m of coaxial cable, 1000 m of other point-to-point link cable, and two repeaters. Each station or repeater connects to the coaxial cable via up to 50 m of “drop cable.” Typical delays associated with each device are given in Table 2.7 (where c = speed of light in a vacuum = 3 × 108 m/s). What is the worst-case round-trip propagation delay, measured in bits, due to the sources listed? (This list is not complete; other sources of delay include sense time and signal rise time.) 42 Coaxial cable Ethernet was limited to a maximum of 500 m between repeaters, which regenerate the signal to 100% of its original amplitude. Along one 500-m segment, the signal could decay to no less than 14% of its original value (8.5 dB). Along 1500 m, then, the decay might be (0.14)3 = 0.3%. Such a signal, even along 2500 m, is still strong enough to be read; why then are repeaters required every 500 m? Item Delay Coaxial cable propagation speed .77c Link/drop cable propagation speed .65c Repeaters approximately 0.6 μs each Transceivers approximately 0.2 μs each Table 2.7 Typical delays associated with various devices (Exercise 41). Exercises 159 43 Suppose the round-trip propagation delay for Ethernet is 46.4 μs. This yields a minimum packet size of 512 bits (464 bits corresponding to propagation delay +48 bits of jam signal). (a) What happens to the minimum packet size if the delay time is held constant, and the signalling rate rises to 100 Mbps? (b) What are the drawbacks to so large a minimum packet size? (c) If compatibility were not an issue, how might the specifications be written so as to permit a smaller minimum packet size? 44 Let A and B be two stations attempting to transmit on an Ethernet. Each has a steady queue of frames ready to send; A’s frames will be numbered A1, A2, and so on, and B’s similarly. Let T = 51.2 μs be the exponential backoff base unit. Suppose A and B simultaneously attempt to send frame 1, collide, and happen to choose backoff times of 0 × T and 1 × T, respectively, meaning A wins the race and transmits A1 while B waits. At the end of this transmission, B will attempt to retransmit B1 while A will attempt to transmit A2. These first attempts will collide, but now A backs off for either 0 × T or 1 × T, while B backs off for time equal to one of 0 × T, ...,3× T. (a) Give the probability that A wins this second backoff race immediately after this first collision; that is, A’s first choice of backoff time k × 51.2 is less than B’s. (b) Suppose A wins this second backoff race. A transmits A3, and when it is finished, A and B collide again as A tries to transmit A4 and B tries once more to transmit B1. Give the probability that A wins this third backoff race immediately after the first collision. (c) Give a reasonable lower bound for the probability that A wins all the remaining backoff races. (d) What then happens to the frame B1? This scenario is known as the Ethernet capture effect. 45 Suppose the Ethernet transmission algorithm is modified as follows: After each suc- cessful transmission attempt, a host waits one or two slot times before attempting to transmit again, and otherwise backs off the usual way. (a) Explain why the capture effect of the previous exercise is now much less likely. (b) Show how the strategy above can now lead to a pair of hosts capturing the Ethernet, alternating transmissions, and locking out a third. 160 2 Direct Link Networks (c) Propose an alternative approach, for example, by modifying the exponential backoff. What aspects of a station’s history might be used as parameters to the modified backoff? 46 Ethernets use Manchester encoding. Assuming that hosts sharing the Ethernet are not perfectly synchronized, why does this allow collisions to be detected soon after they occur, without waiting for the CRC at the end of the packet? 47 Suppose A, B, and C all make their first carrier sense, as part of an attempt to transmit, while a fourth station D is transmitting. Draw a timeline showing one possible sequence of transmissions, attempts, collisions, and exponential backoff choices. Your timeline should also meet the following criteria: (i) initial transmis- sion attempts should be in the order A, B, C, but successful transmissions should be in the order C, B, A, and (ii) there should be at least four collisions. 48 Repeat the previous exercise, now with the assumption that Ethernet is p-persistent with p = 0.33 (that is, a waiting station transmits immediately with probability p when the line goes idle, and otherwise defers one 51.2-μs slot time and repeats the process). Your timeline should meet criterion (i) of the previous problem, but in lieu of criterion (ii), you should show at least one collision and at least one run of four deferrals on an idle line. Again, note that many solutions are possible. 49 Suppose Ethernet physical addresses are chosen at random (using true random bits). (a) What is the probability that on a 1024-host network, two addresses will be the same? (b) What is the probability that the above event will occur on some one or more of 220 networks? (c) What is the probability that of the 230 hosts in all the networks of (b), some pair has the same address? Hint: The calculation for (a) and (c) is a variant of that used in solving the so- called Birthday Problem: Given N people, what is the probability that two of their birthdays (addresses) will be the same? The second person has probability 1 − 1 365 of having a different birthday from the first, the third has probability 1 − 2 365 of having a different birthday from the first two, and so on. The probability all birthdays are different is thus  1 − 1 365 ×  1 − 2 365 ×···× 1 − N − 1 365  Exercises 161 which for smallish N is about 1 − 1 + 2 +···+(N − 1) 365 50 Suppose five stations are waiting for another packet to finish on an Ethernet. All transmit at once when the packet is finished and collide. (a) Simulate this situation up until the point when one of the five waiting stations succeeds. Use coin flips or some other genuine random source to determine backoff times. Make the following simplifications: Ignore interframe spacing, ignore variability in collision times (so that retransmission is always after an exact integral multiple of the 51.2-μs slot time), and assume that each collision uses up exactly one slot time. (b) Discuss the effect of the listed simplifications in your simulation versus the behavior you might encounter on a real Ethernet. 51 Write a program to implement the simulation discussed above, this time with N stations waiting to transmit. Again model time as an integer, T, in units of slot times, and again treat collisions as taking one slot time (so a collision at time T followed by a backoff of k = 0 would result in a retransmission attempt at time T + 1). Find the average delay before one station transmits successfully, for N = 20, N = 40, and N = 100. Does your data support the notion that the delay is linear in N? Hint: For each station, keep track of that station’s NextTimeToSend and CollisionCount. You are done when you reach a time T for which there is only one station with NextTimeToSend == T. If there is no such station, increment T.If there are two or more, schedule the retransmissions and try again. 52 Suppose that N Ethernet stations, all trying to send at the same time, require N/2 slot times to sort out who transmits next. Assuming the average packet size is 5 slot times, express the available bandwidth as a function of N. 53 Consider the following Ethernet model. Transmission attempts are at random times with an average spacing of λ slot times; specifically, the interval between consecutive attempts is an exponential random variable x =−λ log u, where u is chosen randomly in the interval 0 ≤ u ≤ 1. An attempt at time t re- sults in a collision if there is another attempt in the range from t − 1tot + 1, where t is measured in units of the 51.2-μs slot time; otherwise the attempt succeeds. (a) Write a program to simulate, for a given value of λ, the average number of slot times needed before a successful transmission, called the contention interval. 162 2 Direct Link Networks Find the minimum value of the contention interval. Note that you will have to find one attempt past the one that succeeds, in order to determine if there was a collision. Ignore retransmissions, which probably do not fit the random model above. (b) The Ethernet alternates between contention intervals and successful trans- missions. Suppose the average successful transmission lasts 8 slot times (512 bytes). Using your minimum length of the contention interval from above, what fraction of the theoretical 10-Mbps bandwidth is available for transmis- sions? 54 What conditions would have to hold for a corrupted frame to circulate forever on a token ring without a monitor? How does the monitor fix this problem? 55 An IEEE 802.5 token ring has five stations and a total wire length of 230 m. How many bits of delay must the monitor insert into the ring? Do this for both 4 Mbps and 16 Mbps; use a propagation rate of 2.3 × 108 m/s. 56 Consider a token ring network like FDDI in which a station is allowed to hold the token for some period of time (the token holding time, or THT). Let RingLatency denote the time it takes the token to make one complete rotation around the network when none of the stations have any data to send. (a) In terms of THT and RingLatency, express the efficiency of this network when only a single station is active. (b) What setting of THT would be optimal for a network that had only one station active (with data to send) at a time? (c) In the case where N stations are active, give an upper bound on the token rotation time, or TRT, for the network. 57 Consider a token ring with a ring latency of 200 μs. Assuming that the delayed token release strategy is used, what is the effective throughput rate that can be achieved if the ring has a bandwidth of 4 Mbps? What is the effective throughput rate that can be achieved if the ring has a bandwidth of 100 Mbps? Answer for both a single active host and for “many” hosts; for the latter, assume there are sufficiently many hosts transmitting that the time spent advancing the token can be ignored. Assume a packet size of 1 KB. 58 For a 100-Mbps token ring network with a token rotation time of 200 μs that allows each station to transmit one 1-KB packet each time it possesses the token, Exercises 163 calculate the maximum effective throughput rate that any one host can achieve. Do this assuming (a) immediate release and (b) delayed release. 59 Suppose a 100-Mbps delayed-release token ring has 10 stations, a ring latency of 30 μs, and an agreed-upon TTRT of 350 μs. (a) How many synchronous frame bytes could each station send, assuming all are allocated the same amount? (b) Assume stations A, B, C are in increasing order on the ring. Due to uniform synchronous traffic, the TRT without asynchronous data is 300 μs. B sends a 200-μs (2.5-Kb) asynchronous frame. What TRT will A, B, and C then see on their next measurement? Who may transmit such a frame next? Packet Switching Nature seems ...to reach many of her ends by long circuitous routes. —Rudolph Lotze The directly connected networks described in the previous chapter suffer from two limitations. First, there is a limit to how many hosts can be attached. For example, only two hosts can be attached to a point-to-point link, and an Eth- ernet can connect up to only 1024 hosts. Second, there is a limit to how large of a geographic area a single network can serve. For example, an Ethernet can span only PROBLEM Not All Networks Are Directly Connected 2500 m, and even though point-to- point links can be quite long, they do not really serve the area between the two ends. Since our goal is to build networks that can be global in scale, the next problem is therefore to en- able communication between hosts that are not directly connected. This problem is not unlike the one addressed in the telephone network: Your phone is not directly connected to every person you might want to call, but instead is connected to an exchange that contains a switch. It is the switches that create the impression that you have a connection to the person at the other end of the call. Simi- larly, computer networks use packet switches (as distinct from the circuit switches used for telephony) to enable packets to travel from one host to another, even when no direct connection exists between those hosts. This chapter introduces the major concepts of packet switching, which lies at the heart of computer networking. A packet switch is a device with several inputs and outputs leading to and from the hosts that the switch interconnects. The core job of a switch is to take packets that arrive on an input and forward (or switch) them to the right output so that they will reach their appropriate destination. There are a variety of ways that the switch can determine the “right” output for a packet, which can be broadly categorized as connectionless and connection-oriented approaches. 3A key problem that a switch must deal with is the finite bandwidth of its outputs. If packets destined for a certain output arrive at a switch and their arrival rate ex- ceeds the capacity of that output, then we have a problem of contention. The switch queues (buffers) packets until the contention subsides, but if it lasts too long, the switch will run out of buffer space and be forced to discard pack- ets. When packets are discarded too frequently, the switch is said to be congested. The ability of a switch to handle contention is a key aspect of its performance, and many high-performance switches use exotic hardware to reduce the effects of contention. This chapter introduces the issues of forwarding and contention in packet switches. We begin by considering the various approaches to switching, including the connec- tionless and connection-oriented models. We then exam- ine two particular technologies in detail. The first is LAN switching, which has evolved from Ethernet bridging to become one of the dominant technologies in today’s LAN environments. The second noteworthy switching technol- ogy is asynchronous transfer mode (ATM), which is pop- ular among telecommunications service providers in wide area networks. Finally, we consider some of the aspects of switch design that must be taken into account when building large-scale networks. 166 3 Packet Switching 3.1 Switching and Forwarding In the simplest terms, a switch is a mechanism that allows us to interconnect links to form a larger network. A switch is a multi-input, multi-output device, which transfers packets from an input to one or more outputs. Thus, a switch adds the star topol- ogy (see Figure 3.1) to the point-to-point link, bus (Ethernet), and ring (802.5 and FDDI) topologies established in the last chapter. A star topology has several attractive properties: ■ Even though a switch has a fixed number of inputs and outputs, which limits the number of hosts that can be connected to a single switch, large networks can be built by interconnecting a number of switches. ■ We can connect switches to each other and to hosts using point-to-point links, which typically means that we can build networks of large geographic scope. ■ Adding a new host to the network by connecting it to a switch does not necessarily mean that the hosts already connected will get worse performance from the network. This last claim cannot be made for the shared-media networks discussed in the last chapter.For example, it is impossible for two hosts on the same Ethernet to transmit continuously at 10 Mbps because they share the same transmission medium. Every host on a switched network has its own link to the switch, so it may be entirely possible for many hosts to transmit at the full link speed (bandwidth), provided that the switch is designed with enough aggregate capacity. Providing high aggregate throughput is one Figure 3.1 A switch provides a star topology. 3.1 Switching and Forwarding 167 T3 T3 STS-1 Switching protocol Figure 3.2 Example protocol graph running on a switch. Input ports T3 T3 STS-1 T3 T3 STS-1 Switch Output ports Figure 3.3 Example switch with three input and output ports. of the design goals for a switch; we return to this topic below. In general, switched networks are considered more scalable (i.e., more capable of growing to large numbers of nodes) than shared-media networks because of this ability to support many hosts at full speed. A switch is connected to a set of links and, for each of these links, runs the appropriate data link protocol to communicate with the node at the other end of the link. A switch’s primary job is to receive incoming packets on one of its links and to transmit them on some other link. This function is sometimes referred to as either switching or forwarding, and in terms of the OSI architecture, it is the main function of the network layer. Figure 3.2 shows the protocol graph that would run on a switch that is connected to two T3 links and one STS-1 SONET link. A representation of this same switch is given in Figure 3.3. In this figure, we have split the input and output halves of each link, and we refer to each input or output as a port. (In general, we assume that each link is bidirectional, and hence supports both input and output.) In other words, this example switch has three input ports and three output ports. The question then is, How does the switch decide which output port to place each packet on? The general answer is that it looks at the header of the packet for an identifier that it uses to make the decision. The details of how it uses this identifier vary, but there are two common approaches. The first is the datagram or 168 3 Packet Switching connectionless approach. The second is the virtual circuit or connection-oriented ap- proach. A third approach, source routing, is less common than these other two, but it is simple to explain and does have some useful applications. One thing that is common to all networks is that we need to have a way to iden- tify the end nodes. Such identifiers are usually called addresses. We have already seen examples of addresses in the previous chapter, for example, the 48-bit address used for Ethernet. The only requirement for Ethernet addresses is that no two nodes on a network have the same address. This is accomplished by making sure that all Ethernet cards are assigned a globally unique identifier. For the following discussions, we assume that each host has a globally unique address. Later on, we consider other useful prop- erties that an address might have, but global uniqueness is adequate to get us started. Another assumption that we need to make is that there is some way to identify the input and output ports of each switch. There are at least two sensible ways to identify ports: One is to number each port, and the other is to identify the port by the name of the node (switch or host) to which it leads. For now, we use numbering of the ports. 3.1.1 Datagrams The idea behind datagrams is incredibly simple: You just make sure that every packet contains enough information to enable any switch to decide how to get it to its des- tination. That is, every packet contains the complete destination address. Consider the example network illustrated in Figure 3.4, in which the hosts have addresses A, B, C, and so on. To decide how to forward a packet, a switch consults a forwarding table (sometimes called a routing table), an example of which is depicted in Table 3.1. Destination Port A 3 B 0 C 3 D 3 E 2 F 1 G 0 H 0 Table 3.1 Forwarding table for switch 2. 3.1 Switching and Forwarding 169 0 13 2 0 13 2 0 13 2 Switch 3 Host B Switch 2 Host A Switch 1 Host C Host D Host E Host F Host G Host H Figure 3.4 Datagram forwarding: an example network. This particular table shows the forwarding information that switch 2 needs to forward datagrams in the example network. It is pretty easy to figure out such a table when you have a complete map of a simple network like that depicted here; we could imagine a network operator configuring the tables statically. It is a lot harder to create the for- warding tables in large, complex networks with dynamically changing topologies and multiple paths between destinations. That harder problem is known as routing and is the topic of Section 4.2. We can think of routing as a process that takes place in the background so that, when a data packet turns up, we will have the right information in the forwarding table to be able to forward, or switch, the packet. Connectionless (datagram) networks have the following characteristics: ■ A host can send a packet anywhere at any time, since any packet that turns up at a switch can be immediately forwarded (assuming a correctly populated forwarding table). As we will see, this contrasts with most connection-oriented networks, in which some “connection state” needs to be established before the first data packet is sent. ■ When a host sends a packet, it has no way of knowing if the network is capable of delivering it or if the destination host is even up and running. 170 3 Packet Switching ■ Each packet is forwarded independently of previous packets that might have been sent to the same destination. Thus, two successive packets from host A to host B may follow completely different paths (perhaps because of a change in the forwarding table at some switch in the network). ■ A switch or link failure might not have any serious effect on communication if it is possible to find an alternate route around the failure and to update the forwarding table accordingly. This last fact is particularly important to the history of datagram networks. One of the important goals of the ARPANET, forerunner to the Internet, was to develop networking technology that would be robust in a military environment, where you might expect links and nodes to fail because of active attacks such as bombing. It was the ability to route around failures that led to a datagram-based design. 3.1.2 Virtual Circuit Switching A widely used technique for packet switching, which differs significantly from the datagram model, uses the concept of a virtual circuit (VC). This approach, which is also called a connection-oriented model, requires that we first set up a virtual connection from the source host to the destination host before any data is sent. To understand how this works, consider Figure 3.5, where host A again wants to send packets to host B. We can think of this as a two-stage process. The first stage is “connection setup.” The second is data transfer. We consider each in turn. In the connection setup phase, it is necessary to establish “connection state” in each of the switches between the source and destination hosts. The connection state 0 1 2 3 0 1 2 3 0 1 2 3 0 1 2 3 Host A Host B Switch 3 Switch 2Switch 1 Figure 3.5 An example of a virtual circuit network. 3.1 Switching and Forwarding 171 for a single connection consists of an entry in a “VC table” in each switch through which the connection passes. One entry in the VC table on a single switch contains ■ a virtual circuit identifier (VCI) that uniquely identifies the connection at this switch and that will be carried inside the header of the packets that belong to this connection ■ an incoming interface on which packets for this VC arrive at the switch ■ an outgoing interface in which packets for this VC leave the switch ■ a potentially different VCI that will be used for outgoing packets The semantics of one such entry is as follows: If a packet arrives on the designated incoming interface and that packet contains the designated VCI value in its header, then that packet should be sent out the specified outgoing interface with the specified outgoing VCI value first having been placed in its header. Note that the combination of the VCI of packets as they are received at the switch and the interface on which they are received uniquely identifies the virtual connection. There may of course be many virtual connections established in the switch at one time. Also, we observe that the incoming and outgoing VCI values are generally not the same. Thus, the VCI is not a globally significant identifier for the connection; rather, it has significance only on a given link—that is, it has link local scope. Whenever a new connection is created, we need to assign a new VCI for that connection on each link that the connection will traverse. We also need to ensure that the chosen VCI on a given link is not currently in use on that link by some existing connection. There are two broad classes of approach to establishing connection state. One is to have a network administrator configure the state, in which case the virtual circuit is “permanent.” Of course, it can also be deleted by the administrator, so a permanent virtual circuit (PVC) might best be thought of as a long-lived or administratively configured VC. Alternatively, a host can send messages into the network to cause the state to be established. This is referred to as signalling, and the resulting virtual circuits are said to be switched. The salient characteristic of a switched virtual circuit (SVC) is that a host may set up and delete such a VC dynamically without the involvement of a network administrator. Note that an SVC should more accurately be called a “signalled” VC, since it is the use of signalling (not switching) that distinguishes an SVC from a PVC. Let’s assume that a network administrator wants to manually create a new virtual connection from host A to host B. First, the administrator needs to identify a path through the network from A to B. In the example network of Figure 3.5, there is only 172 3 Packet Switching Incoming Outgoing Interface Incoming VCI Interface Outgoing VCI 2 5 1 11 (a) Incoming Outgoing Interface Incoming VCI Interface Outgoing VCI 3 11 2 7 (b) Incoming Outgoing Interface Incoming VCI Interface Outgoing VCI 0 7 1 4 (c) Table 3.2 Virtual circuit table entries for (a) switch 1, (b) switch 2 and (c) switch 3. one such path, but in general this may not be the case. The administrator then picks a VCI value that is currently unused on each link for the connection. For the purposes of our example, let’s suppose that the VCI value 5 is chosen for the link from host A to switch 1, and that 11 is chosen for the link from switch 1 to switch 2. In that case, switch 1 needs to have an entry in its VC table configured as shown in Table 3.2(a). Similarly, suppose that the VCI of 7 is chosen to identify this connection on the link from switch 2 to switch 3, and that a VCI of 4 is chosen for the link from switch 3 to host B. In that case, switches 2 and 3 need to be configured with VC table entries as shown in Table 3.2. Note that the “outgoing” VCI value at one switch is the “incoming” VCI value at the next switch. Once the VC tables have been set up, the data transfer phase can proceed, as illustrated in Figure 3.6. For any packet that it wants to send to host B, A puts the VCI value of 5 in the header of the packet and sends it to switch 1. Switch 1 receives any such packet on interface 2, and it uses the combination of the interface and the VCI in the packet header to find the appropriate VC table entry. As shown in Table 3.2, the table entry in this case tells switch 1 to forward the packet out of interface 1 and 3.1 Switching and Forwarding 173 0 1 2 3 0 13 0 1 2 3 0 1 22 3 Host A Host B Switch 3 Switch 2Switch 1 5 11 Figure 3.6 A packet is sent into a virtual circuit network. 0 1 2 3 0 1 2 3 0 1 2 3 0 1 2 3 Host A Host B Switch 3 Switch 2Switch 1 7 11 Figure 3.7 A packet makes its way through a virtual circuit network. to put the VCI value 11 in the header when the packet is sent. Thus, the packet will arrive at switch 2 on interface 3 bearing VCI 11. Switch 2 looks up interface 3 and VCI 11 in its VC table (as shown in Table 3.2) and sends the packet on to switch 3 after updating the VCI value in the packet header appropriately, as shown in Figure 3.7. This process continues until it arrives at host B with the VCI value of 4 in the packet. To host B, this identifies the packet as having come from host A. In real networks of reasonable size, the burden of configuring VC tables cor- rectly in a large number of switches would quickly become excessive using the above procedures. Thus, some sort of signalling is almost always used, even when setting up “permanent” VCs. In the case of PVCs, signalling is initiated by the network adminis- trator, while SVCs are usually set up using signalling by one of the hosts. We consider now how the same VC just described could be set up by signalling from the host. To start the signalling process, host A sends a setup message into the network, that is, to switch 1. The setup message contains, among other things, the complete destination address of host B. The setup message needs to get all the way to B to create the necessary connection state in every switch along the way. We can see that getting 174 3 Packet Switching the setup message to B is a lot like getting a datagram to B, in that the switches have to know which output to send the setup message to so that it eventually reaches B. For now, let’s just assume that the switches know enough about the network topology to figure out how to do that, so that the setup message flows on to switches 2 and 3 before finally reaching host B. When switch 1 receives the connection request, in addition to sending it on to switch 2, it creates a new entry in its virtual circuit table for this new connection. This entry is exactly the same as shown previously in Table 3.2. The main difference is that now the task of assigning an unused VCI value on the interface is performed by the switch. In this example, the switch picks the value 5. The virtual circuit table now has the following information: “When packets arrive on port 2 with identifier 5, send them out on port 1.” Another issue is that, somehow, host A will need to learn that it should put the VCI value of 5 in packets that it wants to send to B; we will see how that happens below. When switch 2 receives the setup message, it performs a similar process; in this example it picks the value 11 as the incoming VCI value. Similarly, switch 3 picks 7 as the value for its incoming VCI. Each switch can pick any number it likes, as long as that number is not currently in use for some other connection on that port of that switch. As noted above, VCIs have “link local scope”; that is, they have no global significance. Finally, the setup message arrives at host B. Assuming that B is healthy and willing to accept a connection from host A, it too allocates an incoming VCI value, in this case 4. This VCI value can be used by B to identify all packets coming from host A. Now, to complete the connection, everyone needs to be told what their down- stream neighbor is using as the VCI for this connection. Host B sends an acknowledg- ment of the connection setup to switch 3 and includes in that message the VCI that it chose (4). Now switch 3 can complete the virtual circuit table entry for this connection, since it knows the outgoing value must be 4. Switch 3 sends the acknowledgment on to switch 2, specifying a VCI of 7. Switch 2 sends the message on to switch 1, specifying a VCI of 11. Finally, switch 1 passes the acknowledgment on to host A, telling it to use the VCI of 5 for this connection. At this point, everyone knows all that is necessary to allow traffic to flow from host A to host B. Each switch has a complete virtual circuit table entry for the con- nection. Furthermore, host A has a firm acknowledgment that everything is in place all the way to host B. At this point, the connection table entries are in place in all three switches just as in the administratively configured example above, but the whole process happened automatically in response to the signalling message sent from A. The data transfer phase can now begin and is identical to that used in the PVC case. When host A no longer wants to send data to host B, it tears down the connection by sending a teardown message to switch 1. The switch removes the relevant entry from 3.1 Switching and Forwarding 175 its table and forwards the message on to the other switches in the path, which similarly delete the appropriate table entries. At this point, if host A were to send a packet with a VCI of 5 to switch 1, it would be dropped as if the connection had never existed. There are several things to note about virtual circuit switching: ■ Since host A has to wait for the connection request to reach the far side of the network and return before it can send its first data packet, there is at least one RTT of delay before data is sent.1 ■ While the connection request contains the full address for host B (which might be quite large, being a global identifier on the network), each data packet contains only a small identifier, which is only unique on one link. Thus, the per-packet overhead caused by the header is reduced relative to the datagram model. ■ If a switch or a link in a connection fails, the connection is broken and a new one will need to be established. Also, the old one needs to be torn down to free up table storage space in the switches. ■ The issue of how a switch decides which link to forward the connection request on has been glossed over. In essence, this is the same problem as building up the forwarding table for datagram forwarding, which requires some sort of routing algorithm. Routing is described in Section 4.2, and the algorithms described there are generally applicable to routing setup requests as well as datagrams. One of the nice aspects of virtual circuits is that by the time the host gets the go-ahead to send data, it knows quite a lot about the network—for example, that there really is a route to the receiver and that the receiver is willing and able to receive data. It is also possible to allocate resources to the virtual circuit at the time it is established. For example, an X.25 network—a packet-switched network that uses the connection-oriented model—employs the following three-part strategy: 1 Buffers are allocated to each virtual circuit when the circuit is initialized. 2 The sliding window protocol is run between each pair of nodes along the virtual circuit, and this protocol is augmented with flow control to keep the sending node from overrunning the buffers allocated at the receiving node. 3 The circuit is rejected by a given node if not enough buffers are available at that node when the connection request message is processed. 1This is not strictly true. Some people have proposed “optimistically” sending a data packet immediately after sending the connection request. However, most current implementations wait for connection setup to complete before sending data. 176 3 Packet Switching In doing these three things, each node is ensured of having the buffers it needs to queue the packets that arrive on that circuit. This basic strategy is usually called hop-by-hop flow control. By comparison, a datagram network has no connection establishment phase, and each switch processes each packet independently, making it less obvious how a datagram network would allocate resources in a meaningful way. Instead, each arriving packet competes with all other packets for buffer space. If there are no free buffers, the incoming packet must be dis- carded. We observe, however, that even in a datagram-based network, a source host often sends a sequence of packets to the same destination host. It is possible for each switch to distinguish among the set of pack- ets it currently has queued, based on the source/destination pair, and thus for the switch to ensure that the packets belonging to each source/destination pair are receiv- ing a fair share of the switch’s buffers. We discuss this idea in much greater depth in Chapter 6. In the virtual circuit model, we could imagine providing each circuit with a differ- ent quality of service (QoS). In this setting, the term “quality of service” is usually taken to mean that the network gives the user some kind of performance-related guaran- tee, which in turn implies that switches set aside the resources they need to meet this guarantee. For example, the switches along a given virtual circuit might allocate a per- centage of each outgoing link’s bandwidth to that circuit. As another example, a se- quence of switches might ensure that pack- ets belonging to a particular circuit not be delayed (queued) for more than a certain amount of time. We return to the topic of quality of service in Section 6.5. The most popular examples of virtual circuit technologies are Frame Relay and Introduction to Congestion Recall the distinction between con- tention and congestion: Contention occurs when multiple packets have to be queued at a switch be- cause they are competing for the same output link, while congestion means that the switch has so many packets queued that it runs out of buffer space and has to start drop- ping packets. We return to the topic of congestion in Chapter 6, after we have seen the transport proto- col component of the network ar- chitecture. At this point, however, we observe that the decision as to whether your network uses virtual circuits or datagrams has an impact on how you deal with congestion. On the one hand, suppose that each switch allocates enough buffers to handle the packets belonging to each virtual circuit it supports, as is done in an X.25 network. In this case, the net- work has defined away the problem of congestion—a switch never en- counters a situation in which it has more packets to queue than it has buffer space, since it does not allow the connection to be established in the first place unless it can dedicate 3.1 Switching and Forwarding 177 Variable Control 8 Address 16 Frame checksum 16 Flag (0x7E) 8 Flag (0x7E) 8 Data Figure 3.8 Frame Relay packet format. enough resources to it to avoid this situation. The problem with this approach, however, is that it is ex- tremely conservative—it is unlikely that all the circuits will need to use all of their buffers at the same time, and as a consequence, the switch is potentially underutilized. On the other hand, the data- gram model seemingly invites congestion—you do not know that there is enough contention at a switch to cause congestion until you run out of buffers. At that point, it is too late to prevent the congestion, and your only choice is to try to recover from it. The good news, of course, is that you may be able to get better utilization out of your switches since you are not holding buffers in reserve for a worst-case scenario that is unlikely to happen. As is quite often the case, noth- ing is strictly black and white— there are design advantages for defining congestion away (as the X.25 model does) and for doing nothing about congestion until after it happens (as the simple datagram model does). We describe some of these design points in Chapter 6. asynchronous transfer mode (ATM). ATM has a number of interesting properties that we discuss in Section 3.3. Frame Relay is a rather straightforward implementation of virtual circuit technology, and its simplic- ity has made it extremely popular. Many network service providers offer Frame Re- lay PVC services. One of the applications of Frame Relay is the construction of vir- tual private networks (VPNs), a subject dis- cussed in Section 4.1.8. Frame Relay provides some basic quality of service and congestion-avoidance features, but these are rather lightweight compared to X.25 and ATM. The Frame Relay packet format (see Figure 3.8) pro- vides a good example of a packet used for virtual circuit switching. 3.1.3 Source Routing A third approach to switching that uses nei- ther virtual circuits nor conventional data- grams is known as source routing. The name derives from the fact that all the information about network topology that is required to switch a packet across the network is pro- vided by the source host. There are various ways to implement source routing. One would be to assign a number to each output of each switch and to place that number in the header of the packet. The switching function is then very simple: For each packet that arrives on an input, the switch would read the port 178 3 Packet Switching 0 13 2 0 1 3 2 0 13 2 0 13 2 3 0 1 3 01 30 1 Switch 3 Host B Switch 2 Host A Switch 1 Figure 3.9 Source routing in a switched network (where the switch reads the right- most number). number in the header and transmit the packet on that output. However, since there will in general be more than one switch in the path between the sending and the receiving host, the header for the packet needs to contain enough information to allow every switch in the path to determine which output the packet needs to be placed on. One way to do this would be to put an ordered list of switch ports in the header and to rotate the list so that the next switch in the path is always at the front of the list. Figure 3.9 illustrates this idea. In this example, the packet needs to traverse three switches to get from host A to host B. At switch 1, it needs to exit on port 1, at the next switch it needs to exit at port 0, and at the third switch it needs to exit at port 3. Thus, the original header when the packet leaves host A contains the list of ports (3, 0, 1), where we assume that each switch reads the rightmost element of the list. To make sure that the next switch gets the appropriate information, each switch rotates the list after it has read its own entry. Thus, the packet header as it leaves switch 1 en route to switch 2 is now (1, 3, 0); switch 2 performs another rotation and sends out a packet with (0, 1, 3) in the header. Although not shown, switch 3 performs yet another rotation, restoring the header to what it was when host A sent it. There are several things to note about this approach. First, it assumes that host A knows enough about the topology of the network to form a header that has all the right directions in it for every switch in the path. This is somewhat analogous to the problem of building the forwarding tables in a datagram network or figuring out 3.1 Switching and Forwarding 179 Optical Switching To a casual observer of the net- working industry around the year 2000, it might have appeared that the most interesting sort of switch- ing was optical switching. Indeed, optical switching did become an important technology in the late 1990s, due to a confluence of sev- eral factors. One factor was the commercial availability of dense wavelength division multiplexing (DWDM) equipment, which makes it possible to send a great deal of information down a single fiber by transmitting on a large number of optical wavelengths (or colors) at once. Thus, for example, you might send data on 100 or more different wavelengths, and each wavelength might carry as much as 10 Gbps of data. A second factor was the com- mercial availability of optical am- plifiers. Optical signals are atten- uated as they pass through fiber, and after some distance (about 40 km or so) they need to be made stronger in some way. Before opti- cal amplifiers, it was necessary to place repeaters in the path to re- cover the optical signal, convert it to a digital electronic signal, and then convert it back to optical again. Before you could get the data into a repeater, you would have to where to send a setup packet in a virtual cir- cuit network. Second, observe that we can- not predict how big the header needs to be, since it must be able to hold one word of information for every switch on the path. This implies that headers are probably of variable length with no upper bound, un- less we can predict with absolute certainty the maximum number of switches through which a packet will ever need to pass. Third, there are some variations on this approach. For example, rather than rotate the header, each switch could just strip the first element as it uses it. Rotation has an advantage over stripping, however: Host B gets a copy of the complete header, which may help it fig- ure out how to get back to host A. Yet an- other alternative is to have the header carry a pointer to the current “next port” entry, so that each switch just updates the pointer rather than rotating the header; this may be more efficient to implement. We show these three approaches in Figure 3.10. In each case, the entry that this switch needs to read is A, and the entry that the next switch needs to read is B. Source routing can be used in both datagram networks and virtual circuit net- works. For example, the Internet Proto- col, which is a datagram protocol, includes a source route option that allows selected packets to be source routed, while the ma- jority are switched as conventional data- grams. Source routing is also used in some virtual circuit networks as the means to get the initial setup request along the path from source to destination. Finally, we note that source routing suffers from a scaling problem. In any 180 3 Packet Switching Header entering switch Header leaving switch (a) (b) (c) DCBA DC BA DC BA DC B Ptr D C B A Ptr D C B A Figure 3.10 Three ways to handle headers for source routing: (a) rotation; (b) stripping; (c) pointer. The labels are read right to left. reasonably large network, it is very hard for a host to get the complete path information it needs to construct correct headers. 3.2 Bridges and LAN Switches Having discussed some of the basic ideas be- hind switching, we now focus more closely on some specific switching technologies. We begin by considering a class of switches that is used to forward packets between shared-media LANs such as Ethernets. Such switches are sometimes known by the obvi- ous name of LAN switches; historically they have also been referred to as bridges. Suppose you have a pair of Ether- nets that you want to interconnect. One ap- proach you might try is to put a repeater between them, as described in Chapter 2. This would not be a workable solution, however, if doing so exceeded the physi- cal limitations of the Ethernet. (Recall that no more than two repeaters between any pair of hosts and no more than a total of 2500 m in length is allowed.) An alternative demultiplex it using a DWDM ter- minal. Thus, a large number of DWDM terminals would be needed just to drive a single fiber pair for a long distance. Optical amplifiers, unlike repeaters, are analog devices that boost whatever signal is sent along the fiber, even if it is sent on a hundred different wavelengths. Optical amplifiers therefore made DWDM gear much more attrac- tive, because now a pair of DWDM terminals could talk to each other when separated by a distance of hundreds of kilometers. Further- more, you could even upgrade the DWDM gear at the ends without touching the optical amplifiers in the middle of the path, because they will amplify 100 wavelengths as easily as 50 wavelengths. With DWDM and optical am- plifiers, it became possible to build optical networks of huge capac- ity. But at least one more type 3.2 Bridges and LAN Switches 181 would be to put a node between the two Ethernets and have the node forward frames from one Ethernet to the other. This node would be in promiscuous mode, accepting all frames transmitted on either of the Ethernets, so it could forward them to the other. The node we have just described is typically called a bridge, and a collection of LANs connected by one or more bridges is usually said to form an extended LAN.In their simplest variants, bridges simply accept LAN frames on their inputs and forward them out on all other outputs. This simple strategy was used by early bridges, but has since been refined to make bridges a more effective mechanism for interconnecting a set of LANs. The rest of this section fills in the more interesting details. Note that a bridge meets our definition of a switch from the previous section: a multi-input, multi-output device, which transfers packets from an input to one or of device is needed to make these networks useful—the opti- cal switch. Most so-called opti- cal switches today actually perform their switching function electron- ically, and from an architectural point of view they have more in common with the circuit switches of the telephone network than the packet switches described in this chapter. A typical optical switch has a large number of interfaces that understand SONET framing and is able to cross-connect a SONET channel from an incoming interface to an outgoing interface. Thus, with an optical switch, it be- comes possible to provide SONET channels from point A to point B via point C even if there is no di- rect fiber path from A to B—there just needs to be a path from A to C, a switch at C, and a path from C to B. In this respect, an optical more outputs. And recall that this provides a way to increase the total bandwidth of a network. For example, while a single Ethernet segment can carry only 10 Mbps of total traffic, an Ethernet bridge can carry as much as 10n Mbps, where n is the num- ber of ports (inputs and outputs) on the bridge. 3.2.1 Learning Bridges The first optimization we can make to a bridge is to observe that it need not for- ward all frames that it receives. Consider the bridge in Figure 3.11. Whenever a frame from host A that is addressed to host B arrives on port 1, there is no need for the bridge to forward the frame out over port 2. The question, then, is, How does a bridge come to learn on which port the various hosts reside? One option would be to have a human download a table into the bridge similar to the one given in Table 3.3. Then, whenever the bridge receives a frame on port 1 that is addressed to host A, it would not forward 182 3 Packet Switching the frame out on port 2; there would be no need because host A would have already di- rectly received the frame on the LAN con- nected to port 1. Anytime a frame addressed to host A was received on port 2, the bridge would forward the frame out on port 1. Note that a bridge using such a table would be using the datagram (or connec- tionless) model of forwarding described in Section 3.1.1. Each packet carries a global address, and the bridge decides which out- put to send a packet on by looking up that address in a table. Having a human maintain this table is quite a burden, especially considering that there is a simple trick by which a bridge can learn this information for itself. The idea is for each bridge to inspect the source address in all the frames it receives. Thus, when host A sends a frame to a host on either side of the bridge, the bridge receives this frame and records the fact that a frame from host A was just received on port 1. In this way, the bridge can build a table just like Table 3.3. When a bridge first boots, this table is empty; entries are added over time. Also, a timeout is associated with each entry, and the bridge discards the entry after a speci- fied period of time. This is to protect against the situation in which a host—and as a consequence, its LAN address—is moved from one network to another. Thus, this ta- ble is not necessarily complete. Should the bridge receive a frame that is addressed to a host not currently in the table, it goes ahead and forwards the frame out on all the switch bears some relationship to the switches in Figure 3.5, in that it creates the illusion of a con- netion between two points even when there is no direct physical connection between them. How- ever, optical switches do not pro- vide virtual circuits; they provide “real” circuits (e.g., a SONET channel). There are even some newer types of optical switches that use microscopic mirrors to deflect all the light from one switch port to another, so that there could be an uninterrupted optical channel from point A to point B. We don’t cover optical net- working extensively in this book, in part because of space considera- tions. For many practical purposes, you can think of optical networks as a piece of the infrastructure that enables telephone companies to provide SONET links or other types of circuits where and when you need them. However, it is worth noting that many of the tech- nologies that are discussed later in this book, such as routing protocols and Multiprotocol Label Switch- ing, do have application to the world of optical networking. 3.2 Bridges and LAN Switches 183 A Bridge B C XYZ Port 1 Port 2 Figure 3.11 Illustration of a learning bridge. Host Port A 1 B 1 C 1 X 2 Y 2 Z 2 Table 3.3 Forwarding table maintained by a bridge. other ports. In other words, this table is simply an optimization that filters out some frames; it is not required for correctness. Implementation The code that implements the learning bridge algorithm is quite simple, and we sketch it here. Structure BridgeEntry defines a single entry in the bridge’s forwarding table; these are stored in a Map structure (which supports mapCreate, mapBind, and MapResolve operations) to enable entries to be efficiently located when packets arrive from sources already in the table. The constant MAX TTL specifies how long an entry is kept in the table before it is discarded. 184 3 Packet Switching #define BRIDGE_TAB_SIZE 1024 /* max. size of bridging table */ #define MAX_TTL 120 /* time (in seconds) before an entry is flushed */ typedef struct { MacAddr destination; /* MAC address of a node */ int ifnumber; /* interface to reach it */ u_short TTL; /* time to live */ Binding binding; /* binding in the Map */ } BridgeEntry; int numEntries = 0; Map bridgeMap = mapCreate(BRIDGE_TAB_SIZE, sizeof(BridgeEntry)); The routine that updates the forwarding table when a new packet arrives is given by updateTable. The arguments passed are the source MAC address contained in the packet and the interface number on which it was received. Another routine, not shown here, is invoked at regular intervals, scans the entries in the forwarding table, and decrements the TTL (time to live) field of each entry, discarding any entries whose TTL has reached 0. Note that the TTL is reset to MAX TTL every time a packet arrives to refresh an existing table entry, and that the interface on which the destination can be reached is updated to reflect the most recently received packet. void updateTable (MacAddr src, int inif) { BridgeEntry *b; if (mapResolve(bridgeMap, &src, (void **)&b) == FALSE) { /* this address is not in the table, so try to add it */ if (numEntries < BRIDGE_TAB_SIZE) { b = NEW(BridgeEntry); b->binding = mapBind( bridgeMap, &src, b); /* use source address of packet as dest. address in table */ b->destination = src; numEntries++; } else { /* can't fit this address in the table now, so give up */ 3.2 Bridges and LAN Switches 185 return; } } /* reset TTL and use most recent input interface */ b->TTL = MAX_TTL; b->ifnumber = inif; } Note that this implementation adopts a simple strategy in the case where the bridge table has become full to capacity—it simply fails to add the new address. Recall that completeness of the bridge table is not necessary for correct forwarding; it just optimizes performance. If there is some entry in the table that is not currently being used, it will eventually time out and be removed, creating space for a new entry. An alternative approach would be to invoke some sort of cache replacement algorithm on finding the table full; for example, we might locate and remove the entry with the smallest TTL to accommodate the new entry. 3.2.2 Spanning Tree Algorithm The preceding strategy works just fine until the extended LAN has a loop in it, in which case it fails in a horrible way—frames potentially loop through the extended LAN forever. This is easy to see in the example depicted in Figure 3.12, where, for example, bridges B1, B4, and B6 form a loop. How does an extended LAN come to have a loop in it? One possibility is that the network is managed by more than one administrator, for example, because it spans multiple departments in an organization. In such a setting, it is possible that no single person knows the entire configuration of the network, meaning that a bridge that closes a loop might be added without anyone knowing. A second, more likely scenario is that loops are built into the network on purpose—to provide redundancy in case of failure. Whatever the cause, bridges must be able to correctly handle loops. This problem is addressed by having the bridges run a distributed spanning tree algorithm. If you think of the extended LAN as being represented by a graph that possibly has loops (cycles), then a spanning tree is a subgraph of this graph that covers (spans) all the vertices, but contains no cycles. That is, a spanning tree keeps all of the vertices of the original graph, but throws out some of the edges. For example, Figure 3.13 shows a cyclic graph on the left and one of possibly many spanning trees on the right. The spanning tree algorithm, which was developed by Radia Perlman at Digital, is a protocol used by a set of bridges to agree upon a spanning tree for a particu- lar extended LAN. (The IEEE 802.1 specification for LAN bridges is based on this algorithm.) In practice, this means that each bridge decides the ports over which it 186 3 Packet Switching A C E D B K F H J G I B3 B7 B4 B2 B5 B1 B6 Figure 3.12 Extended LAN with loops. (a) (b) Figure 3.13 Example of (a) a cyclic graph; (b) a corresponding spanning tree. is and is not willing to forward frames. In a sense, it is by removing ports from the topology that the extended LAN is reduced to an acyclic tree.2 It is even possible that an entire bridge will not participate in forwarding frames, which seems strange when you consider that the one reason we intentionally have loops in the network in the first 2Representing an extended LAN as an abstract graph is a bit awkward. Basically, you let both the bridges and the LANs correspond to the vertices of the graph and the ports correspond to the graph’s edges. However, the spanning tree we are going to compute for this graph needs to span only those nodes that correspond to networks. It is possible that nodes corresponding to bridges will be disconnected from the rest of the graph. This corresponds to a situation in which all the ports connecting a bridge to various networks get removed by the algorithm. 3.2 Bridges and LAN Switches 187 place is to provide redundancy. The algorithm is dynamic, however, meaning that the bridges are always prepared to reconfigure themselves into a new spanning tree should some bridge fail. The main idea of the spanning tree is for the bridges to select the ports over which they will forward frames. The algorithm selects ports as follows. Each bridge has a unique identifier; for our purposes, we use the labels B1, B2, B3, and so on. The algorithm first elects the bridge with the smallest id as the root of the spanning tree; exactly how this election takes place is described below. The root bridge always for- wards frames out over all of its ports. Next, each bridge computes the shortest path to the root and notes which of its ports is on this path. This port is also selected as the bridge’s preferred path to the root. Finally, all the bridges connected to a given LAN elect a single designated bridge that will be responsible for forwarding frames toward the root bridge. Each LAN’s designated bridge is the one that is closest to the root, and if two or more bridges are equally close to the root, then the bridges’ identifiers are used to break ties; the smallest id wins. Of course, each bridge is connected to more than one LAN, so it participates in the election of a designated bridge for each LAN it is connected to. In effect, this means that each bridge decides if it is the designated bridge relative to each of its ports. The bridge forwards frames over those ports for which it is the designated bridge. Figure 3.14 shows the spanning tree that corresponds to the extended LAN shown in Figure 3.12. In this example, B1 is the root bridge, since it has the smallest id. Notice that both B3 and B5 are connected to LAN A, but B5 is the designated bridge since it is closer to the root. Similarly, both B5 and B7 are connected to LAN B, but in this case, B5 is the designated bridge since it has the smaller id; both are an equal distance from B1. While it is possible for a human to look at the extended LAN given in Figure 3.12 and to compute the spanning tree given in Figure 3.14 according to the rules given above, the bridges in an extended LAN do not have the luxury of being able to see the topology of the entire network, let alone peek inside other bridges to see their ids. Instead, the bridges have to exchange configuration messages with each other and then decide whether or not they are the root or a designated bridge based on these messages. Specifically, the configuration messages contain three pieces of information: 1 the id for the bridge that is sending the message 2 the id for what the sending bridge believes to be the root bridge 3 the distance, measured in hops, from the sending bridge to the root bridge Each bridge records the current “best” configuration message it has seen on each of its ports (“best” is defined below), including both messages it has received from other bridges and messages that it has itself transmitted. 188 3 Packet Switching A C E D B K F H J G I B5 B2 B3 B7 B4 B1 B6 Figure 3.14 Spanning tree with some ports not selected. Initially, each bridge thinks it is the root, and so it sends a configuration message out on each of its ports identifying itself as the root and giving a distance to the root of 0. Upon receiving a configuration message over a particular port, the bridge checks to see if that new message is better than the current best configuration message recorded for that port. The new configuration message is considered “better” than the currently recorded information if ■ it identifies a root with a smaller id or ■ it identifies a root with an equal id but with a shorter distance or ■ the root id and distance are equal, but the sending bridge has a smaller id. If the new message is better than the currently recorded information, the bridge discards the old information and saves the new information. However, it first adds 1 to the distance-to-root field since the bridge is one hop farther away from the root than the bridge that sent the message. When a bridge receives a configuration message indicating that it is not the root bridge—that is, a message from a bridge with a smaller id—the bridge stops generating configuration messages on its own and instead only forwards configuration messages from other bridges, after first adding 1 to the distance field. Likewise, when a bridge 3.2 Bridges and LAN Switches 189 receives a configuration message that indicates it is not the designated bridge for that port—that is, a message from a bridge that is closer to the root or equally far from the root but with a smaller id—the bridge stops sending configuration messages over that port. Thus, when the system stabilizes, only the root bridge is still generating configuration messages, and the other bridges are forwarding these messages only over ports for which they are the designated bridge. To make this more concrete, consider what would happen in Figure 3.14 if the power had just been restored to the building housing this network, so that all the bridges boot at about the same time. All the bridges would start off by claiming to be the root. We denote a configuration message from node X in which it claims to be distance d from root node Y as (Y, d, X). Focusing on the activity at node B3, a sequence of events would unfold as follows: 1 B3 receives (B2, 0, B2). 2 Since 2 < 3, B3 accepts B2 as root. 3 B3 adds one to the distance advertised by B2 (0) and thus sends (B2, 1, B3) toward B5. 4 Meanwhile, B2 accepts B1 as root because it has the lower id, and it sends (B1, 1, B2) toward B3. 5 B5 accepts B1 as root and sends (B1, 1, B5) toward B3. 6 B3 accepts B1 as root, and it notes that both B2 and B5 are closer to the root than it is. Thus B3 stops forwarding messages on both its interfaces. This leaves B3 with both ports not selected, as shown in Figure 3.14. Even after the system has stabilized, the root bridge continues to send configura- tion messages periodically, and the other bridges continue to forward these messages as described in the previous paragraph. Should a particular bridge fail, the downstream bridges will not receive these configuration messages, and after waiting a specified period of time, they will once again claim to be the root, and the algorithm just des- cribed will kick in again to elect a new root and new designated bridges. One important thing to notice is that although the algorithm is able to recon- figure the spanning tree whenever a bridge fails, it is not able to forward frames over alternative paths for the sake of routing around a congested bridge. 3.2.3 Broadcast and Multicast The preceding discussion has focused on how bridges forward unicast frames from one LAN to another. Since the goal of a bridge is to transparently extend a LAN across multiple networks, and since most LANs support both broadcast and multicast, 190 3 Packet Switching then bridges must also support these two features. Broadcast is simple—each bridge forwards a frame with a destination broadcast address out on each active (selected) port other than the one on which the frame was received. Multicast can be implemented in exactly the same way, with each host deciding for itself whether or not to accept the message. This is exactly what is done in practice. Notice, however, that since not all the LANs in an extended LAN necessarily have a host that is a member of a particular multicast group, it is possible to do better. Specifically, the spanning tree algorithm can be extended to prune networks over which multicast frames need not be forwarded. Consider a frame sent to group M by a host on LAN A in Figure 3.14. If there is no host on LAN J that belongs to group M, then there is no need for bridge B4 to forward the frames over that network. On the other hand, not having a host on LAN H that belongs to group M does not necessarily mean that bridge B1 can avoid forwarding multicast frames onto LAN H. It all depends on whether or not there are members of group M on LANs I and J. How does a given bridge learn whether it should forward a multicast frame over a given port? It learns exactly the same way that a bridge learns whether it should for- ward a unicast frame over a particular port—by observing the source addresses that it receives over that port. Of course, groups are not typically the source of frames, so we have to cheat a little. In particular, each host that is a member of group M must period- ically send a frame with the address for group M in the source field of the frame header. This frame would have as its destination address the multicast address for the bridges. Note that while the multicast extension just described has been proposed, it is not widely adopted. Instead, multicast is implemented in exactly the same way as broadcast on today’s extended LANs. 3.2.4 Limitations of Bridges The bridge-based solution just described is meant to be used in only a fairly limited setting—to connect a handful of similar LANs. The main limitations of bridges become apparent when we consider the issues of scale and heterogeneity. On the issue of scale, it is not realistic to connect more than a few LANs by means of bridges, where in practice “few” typically means “tens of.” One reason for this is that the spanning tree algorithm scales linearly; that is, there is no provision for imposing a hierarchy on the extended LAN. A second reason is that bridges forward all broadcast frames. While it is reasonable for all hosts within a limited setting (say, a department) to see each other’s broadcast messages, it is unlikely that all the hosts in a larger environment (say, a large company or university) would want to have to be bothered by each other’s broadcast messages. Said another way, broadcast does not scale, and as a consequence, extended LANs do not scale. 3.2 Bridges and LAN Switches 191 WX B1 B2 Y Z VLAN 100 VLAN 100 VLAN 200 VLAN 200 Figure 3.15 Two virtual LANs share a common backbone. One approach to increasing the scalability of extended LANs is the virtual LAN (VLAN). VLANs allow a single extended LAN to be partitioned into several seemingly separate LANs. Each virtual LAN is assigned an identifier (sometimes called a color), and packets can only travel from one segment to another if both segments have the same identifier. This has the effect of limiting the number of segments in an extended LAN that will receive any given broadcast packet. We can see how VLANs work with an example. Figure 3.15 shows four hosts on four different LAN segments. In the absence of VLANs, any broadcast packet from any host will reach all the other hosts. Now let’s suppose that we define the segments connected to hosts W and X as being in one VLAN, which we’ll call VLAN 100. We also define the segments that connect to hosts Y and Z as being in VLAN 200. To do this, we need to configure a VLAN ID on each port of bridges B1 and B2. The link between B1 and B2 is considered to be in both VLANs. When a packet sent by host X arrives at bridge B2, the bridge observes that it came in a port that was configured as being in VLAN 100. It inserts a VLAN header between the Ethernet header and its payload. The interesting part of the VLAN header is the VLAN ID; in this case, that ID is set to 100. The bridge now applies its normal rules for forwarding to the packet, with the extra restriction that the packet may not be sent out an interface that is not part of VLAN 100. Thus, under no circumstances will the packet—even a broadcast packet—be sent out the interface to host Z, which is in VLAN 200. The packet is, however, forwarded to bridge B1, which follows the same rules, and thus may forward the packet to host W but not to host Y. An attractive feature of VLANs is that it is possible to change the logical topology without moving any wires or changing any addresses. For example, if we wanted to make the segment that connects to host Z be part of VLAN 100, and thus enable X, W, and Z to be on the same virtual LAN, we would just need to change one piece of configuration on bridge B2. 192 3 Packet Switching On the issue of heterogeneity, bridges are fairly limited in the kinds of networks they can interconnect. In particular, bridges make use of the network’s frame header and so can support only networks that have exactly the same format for addresses. Thus, bridges can be used to connect Ethernets to Ethernets, 802.5 to 802.5, and Ethernets to 802.5 rings, since both networks support the same 48-bit address format. Bridges do not readily generalize to other kinds of networks, such as ATM.3 Despite their limitations, bridges are a very important part of the complete net- working picture. Their main advantage is that they allow multiple LANs to be trans- parently connected; that is, the networks can be connected without the end hosts having to run any additional protocols (or even be aware, for that matter). The one potential exception is when the hosts are expected to announce their membership in a multicast group, as described in Section 3.2.3. Notice, however, that this transparency can be dangerous. If a host, or more pre- cisely, the application and transport protocol running on that host, is programmed under the assumption that it is running on a single LAN, then inserting bridges between the source and destination hosts can have unexpected consequences. For example, if a bridge becomes congested, it may have to drop frames; in contrast, it is rare that a single Ethernet ever drops a frame. As another example, the latency between any pair of hosts on an extended LAN becomes both larger and more highly variable; in contrast, the physical limitations of a single Ethernet make the latency both small and predictable. As a final example, it is possible (although unlikely) that frames will be reordered in an extended LAN; in contrast, frame order is never shuffled on a single Ethernet. The bottom line is that it is never safe to design network software under the assumption that it will run over a single Ethernet segment. Bridges happen. 3.3 Cell Switching (ATM) Another switching technology that deserves special attention is asynchronous transfer mode (ATM). ATM became an important technology in the 1980s and early 1990s for a variety of reasons, not the least of which is that it was embraced by the telephone industry, which has historically been less than active in data communications except as a supplier of links on top of which other people have built networks. ATM also happened to be in the right place at the right time, as a high-speed switching technology that appeared on the scene just when shared media like Ethernet and 802.5 were starting to look a bit too slow for many users of computer networks. In some ways ATM is a competing technology with Ethernet switching, but the areas of application for these two technologies only partially overlap. 3As we will see in Section 3.3.5, there are techniques to make ATM networks look more like “conventional” LANs such as Ethernets, and bridges do have a role in this environment. 3.3 Cell Switching (ATM) 193 ATM is a connection-oriented, packet-switched technology, which is to say, it uses virtual circuits very much in the manner described in Section 3.1.2. In ATM terminology, the connection setup phase is called signalling. The main ATM signalling protocol is known as Q.2931. In addition to discovering a suitable route across an ATM network, Q.2931 is also responsible for allocating resources at the switches along the circuit. This is done in an effort to ensure the circuit a particular quality of service. Indeed, the QoS capabilities of ATM are one of its greatest strengths. We return to this topic in Chapter 6, where we discuss it in the context of similar efforts to implement QoS. When any virtual connection is set up, it is necessary to put the address of the destination in the signalling message. In ATM, this address can be in one of several formats, the most common ones being E.164 and NSAP (network service access point); the details are not terribly important here, except to note that they are different from the MAC addresses used in traditional LANs. One thing that makes ATM really unusual is that the packets that are switched in an ATM network are of fixed length. That length happens to be 53 bytes—5 bytes of header followed by 48 bytes of payload—a rather interesting choice that is discussed in more detail below. To distinguish these fixed-length packets from the more common variable-length packets normally used in computer networks, they are given a special name: cells. ATM may be thought of as the canonical example of cell switching. 3.3.1 Cells All the packet-switching technologies we have looked at so far have used variable- length packets. Variable-length packets are normally constrained to fall within some bounds. The lower bound is set by the minimum amount of information that needs to be contained in the packet, which is typically a header with no optional extensions. The upper bound may be set by a variety of factors; the maximum FDDI packet size, for example, determines how long each station is allowed to transmit without passing on the token, and thus determines how long a station might have to wait for the token to reach it. Cells, in contrast, are both fixed in length and small in size. While this seems like a simple enough design choice, there are actually a lot of factors involved, as explained in the following paragraphs. Cell Size Variable-length packets have some nice characteristics. If you only have 1 byte to send (e.g., to acknowledge the receipt of a packet), you put it in a minimum-sized packet. If you have a large file to send, however, you break it up into as many maximum- sized packets as you need. You do not need to send any extraneous padding in the first case, and in the second, you drive down the ratio of header to data bytes, thus 194 3 Packet Switching increasing bandwidth efficiency. You also minimize the total number of packets sent, thereby minimizing the total processing incurred by per-packet operations. This can be particularly important in obtaining high throughput, since many network devices are limited not by how many bits per second they can process but rather by the number of packets per second. So, why use fixed-length cells? One of the main reasons was to facilitate the implementation of hardware switches. When ATM was being created in the mid- and late 1980s, 10-Mbps Ethernet was the cutting-edge technology in terms of speed. To go much faster, most people thought in terms of hardware. Also, in the telephone world, people think big when they think of switches—telephone switches often serve tens of thousands of customers. Fixed-length packets turn out to be a very helpful thing if you want to build fast, highly scalable switches. There are two main reasons for this: 1 It is easier to build hardware to do simple jobs, and the job of processing packets is simpler when you already know how long each one will be. 2 If all packets are the same length, then you can have lots of switching elements all doing much the same thing in parallel, each of them taking the same time to do its job. This second reason, the enabling of parallelism, greatly improves the scalability of switch designs. It would be overstating the case to say that fast parallel hardware switches can only be built using fixed-length cells. However, it is certainly true that cells ease the task of building such hardware and that there was a lot of knowledge available about how to build cell switches in hardware at the time the ATM standards were being defined. Another nice property of cells relates to the behavior of queues. Queues build up in a switch when traffic from several inputs may be heading for a single output. In general, once you extract a packet from a queue and start transmitting it, you need to continue until the whole packet is transmitted; it is not practical to preempt the transmission of a packet. The longest time that a queue output can be tied up is equal to the time it takes to transmit a maximum-sized packet. Fixed-length cells mean that a queue output is never tied up for more than the time it takes to transmit one cell, which is almost certainly shorter than the maximum-sized packet on a variable-length packet network. Thus, if tight control over the latency that is being experienced by cells when they pass through a queue is important, cells provide some advantage. Of course, long queues can still build up, and there is no getting around the fact that some cells will have to wait their turn. What you get from cells is not much shorter queues but potentially finer control over the behavior of queues. 3.3 Cell Switching (ATM) 195 An example will help to clarify this idea. Imagine a network with variable-length packets, where the maximum packet length is 4 KB and the link speed is 100 Mbps. The time to transmit a maximum-sized packet is 4096 × 8/100 = 327.68 μs. Thus, a high-priority packet that arrives just after the switch starts to transmit a 4-KB packet will have to sit in the queue 327.68 μs waiting for access to the link. In contrast, if the switch were forwarding 53-byte cells, the longest wait would be 53×8/100 = 4.24 μs. This may not seem like a big deal, but the ability to control delay and especially to control its variation with time (jitter) can be important for some applications. Queues of cells also tend to be a little shorter than queues of packets, for the following reason. When a packet begins to arrive in an empty queue, it is typical for the switch to have to wait for the whole packet to arrive before it can start transmitting the packet on an outgoing link. This means that the link sits idle while the packet arrives. However, if you imagine a large packet being replaced by a “train” of small cells, then as soon as the first cell in the train has entered the queue, the switch can transmit it. Imagine in the example above what would happen if two 4-KB packets arrived in a queue at about the same time. The link would sit idle for 327.68 μs while these two packets arrive, and at the end of that period we would have 8 KB in the queue. Only then could the queue start to empty. If those same two packets were sent as trains of cells, then transmission of the cells could start 4.24 μs after the first train started to arrive. At the end of 327.68 μs, the link would have been active for a little over 323 μs, and there would be just over 4 KB of data left in the queue, not 8 KB as before. Shorter queues mean less delay for all the traffic. Having decided to use small, fixed-length packets, the next question is, What is the right length to fix them at? If you make them too short, then the amount of header information that needs to be carried around relative to the amount of data that fits in one cell gets larger, so the percentage of link bandwidth that is actually used to carry data goes down. Even more seriously, if you build a device that processes cells at some maximum number of cells per second, then as cells get shorter, the total data rate drops in direct proportion to cell size. An example of such a device might be a network adaptor that reassembles cells into larger units before handing them up to the host. The performance of such a device depends directly on cell size. On the other hand, if you make the cells too big, then there is a problem of wasted bandwidth caused by the need to pad transmitted data to fill a complete cell. If the cell payload size is 48 bytes and you want to send 1 byte, you’ll need to send 47 bytes of padding. If this happens a lot, then the utilization of the link will be very low. Efficient link utilization is not the only factor that influences cell size. For ex- ample, cell size has a particular effect on voice traffic, and since ATM grew out of the telephony community, one of the major concerns was that it be able to carry 196 3 Packet Switching voice effectively. The standard digital en- coding of voice is done at 64 Kbps (8-bit samples taken at 8 KHz). To maximize ef- ficiency, you want to collect a full cell’s worth of voice samples before transmitting a cell. A sampling rate of 8 KHz means that 1 byte is sampled every 125 μs, so the time it takes to fill an n-byte cell with samples is n × 125 μs. If cells are, say, 1000 bytes long, it would take 125 ms just to collect a full cell of samples before you even start to transmit it to the receiver. That amount of latency starts to be quite noticeable to a human listener. Even considerably smaller latencies create problems for voice, partic- ularly in the form of echoes. Echoes can be eliminated by a piece of technology called an echo canceler, but this adds cost to a tele- phone network that many network opera- tors would rather avoid. All of the above factors caused a great deal of debate in the international standards bodies when ATM was being standardized, and the fact that no length was perfect in all cases was used by those opposed to ATM to argue that fixed-length cells were a bad idea in the first place. As is so often the case with standards, the end result was a compromise that pleased almost no one: 48 bytes was chosen as the length for the ATM cell pay- load. Probably the greatest tragedy of this choice is that it is not a power of two, which means that it is quite a mismatch to most things that computers handle, like pages and cache lines. Rather less controversially, the header was fixed at 5 bytes. The format of an ATM cell is shown in Figure 3.16; note that this figure shows the field lengths in bits. A Compromise of 48 Bytes The explanation for why the pay- load of an ATM cell is 48 bytes is an interesting one and is an ex- cellent case study in the process of standardization. As the ATM standard was evolving, the U.S. telephone companies were pushing for a 64-byte cell size, while the European companies were advo- cating 32-byte cells. The reason that the Europeans wanted the smaller size was that since the coun- tries they served were of a small enough size, they would not have to install echo cancelers if they were able to keep the latency induced by generating a complete cell small enough. Thirty-two-byte cells were adequate for this purpose. In con- trast, the United States is a large enough country that the phone companies had to install echo can- celers anyway, and so the larger cell size reflected a desire to improve the header-to-payload ratio. Averaging is a classic form of compromise—48 bytes is sim- ply the average of 64 bytes and 32 bytes. So as not to leave the false impression that this use of compromise-by-averaging is an iso- lated incident, we note that the seven-layer OSI model was actu- ally a compromise between six and eight layers. 3.3 Cell Switching (ATM) 197 GFC HEC (CRC-8) 416318 VPI VCI CLPType Payload 384 (48 bytes)8 Figure 3.16 ATM cell format at the UNI. Cell Format The ATM cell actually comes in two different formats, depending on where you look in the network. The one shown in Figure 3.16 is called the UNI (user-network inter- face) format; the alternative is the NNI (network-network interface). The UNI format is used, of course, at the user-to-network interface. This is likely to be the interface between a telephone company and one of its customers. The network-to-network interface is likely to be between a pair of phone companies. The only significant dif- ference in cell formats is that the NNI format replaces the GFC field with 4 extra bits of VPI. Clearly, understanding all the three-letter acronyms (TLAs) is a key part of understanding ATM. Starting from the leftmost byte of the cell (which is the first one transmitted), the UNI cell has 4 bits for generic flow control (GFC). These bits have not been widely used, but they were intended to have local significance at a site and could be overwritten in the network. The basic idea behind the GFC bits was to provide a means to arbitrate access to the link if the local site used some shared medium to connect to ATM. The next 24 bits contain an 8-bit virtual path identifier (VPI) and a 16-bit virtual circuit identifier (VCI). The difference between the two is explained below, but for now it is adequate to think of them as a single 24-bit identifier that is used to identify a virtual connection, just as in Section 3.1.2. Following the VPI/VCI is a 3-bit Type field that has eight possible values. Four of them, when the first bit in the field is set, relate to management functions. When that bit is clear, it means that the cell contains user data. In this case, the second bit is the “explicit forward congestion indication” (EFCI) bit, and the third is the “user signalling” bit. The former can be set by a congested switch to tell an end node that it is congested; it has its roots in the DECbit described in Section 6.4.1; in ATM, it is used for congestion control in conjunction with the available bit rate (ABR) service class described in Section 6.5.4. The third bit is used primarily in conjunction with ATM Adaptation Layer 5 to delineate frames, as discussed below. Next is a bit to indicate cell loss priority (CLP); a user or network element may set this bit to indicate cells that should be dropped preferentially in the event of overload. For example, a video coding application could set this bit for cells that, if dropped, would not dramatically degrade the quality of the video. A network element might set 198 3 Packet Switching this bit for cells that have been transmitted by a user in excess of the amount that was negotiated. The last byte of the header is an 8-bit CRC, known as the header error check (HEC). It uses the CRC-8 polynomial given in Section 2.4.3 and provides error detection and single-bit error correction capability on the cell header only. Protecting the cell header is particularly important because an error in the VCI will cause the cell to be misdelivered. 3.3.2 Segmentation and Reassembly Up to this point, we have assumed that a low-level protocol could just accept the packet handed down to it by a high-level protocol, attach its own header, and pass the packet on down. This is not possible with ATM, however, since the packets handed down from above are often larger than 48 bytes, and thus, will not fit in the payload of an ATM cell. The solution to this problem is to fragment the high-level message into low-level packets at the source, transmit the individual low-level packets over the network, and then reassemble the fragments back together at the destination. This general technique is usually called fragmentation and reassembly. In the case of ATM, however, it is often called segmentation and reassembly (SAR). Segmentation is not unique to ATM, but it is much more of a problem than in a network with a maximum packet size of, say, 1500 bytes. To address the issue, a pro- tocol layer was added that sits between ATM and the variable-length packet protocols that might use ATM, such as IP. This layer is called the ATM Adaptation Layer (AAL), and to a first approximation, the AAL header simply contains the information needed by the destination to reassemble the individual cells back into the original message. The relationship between the AAL and ATM is illustrated in Figure 3.17. Because ATM was designed to support all sorts of services, including voice, video, and data, it was felt that different services would have different AAL needs. …… AAL ATM AAL ATM Figure 3.17 Segmentation and reassembly in ATM. 3.3 Cell Switching (ATM) 199 Thus, four adaptation layers were originally defined: 1 and 2 were designed to support applications, like voice, that require guaranteed bit rates, while 3 and 4 were intended to provide support for packet data running over ATM. The idea was that AAL3 would be used by connection-oriented packet services (such as X.25) and AAL4 would be used by connectionless services (such as IP). Eventually, the reasons for having different AALs for these two types of service were found to be insufficient, and the AALs merged into one that is inconveniently known as AAL3/4. Meanwhile, some perceived shortcomings in AAL3/4 caused a fifth AAL to be proposed, called AAL5. Thus, there are now four AALs: 1, 2, 3/4, and 5. The two that support computer communications are described below. ATM Adaptation Layer 3/4 The main function of AAL3/4 is to provide enough information to allow variable- length packets to be transported across the ATM network as a series of fixed-length cells. That is, the AAL supports the segmentation and reassembly process. Since we are now working at a new layer of the network hierarchy, convention requires us to introduce a new name for a packet—in this case, we call it a protocol data unit (PDU). The task of segmentation/reassembly involves two different packet formats. The first of these is the convergence sublayer protocol data unit (CS-PDU), as depicted in Figure 3.18. The CS-PDU defines a way of encapsulating variable-length PDUs prior to segmenting them into cells. The PDU passed down to the AAL layer is encapsulated by adding a header and a trailer, and the resultant CS-PDU is segmented into ATM cells. The CS-PDU format begins with an 8-bit common part indicator (CPI), which indicates which version of the CS-PDU format is in use. Only the value 0 is currently defined. The next 8 bits contain the beginning tag (Btag), which is supposed to match the end tag (Etag) for a given PDU. This protects against the situation in which the loss of the last cell of one PDU and the first cell of another causes two PDUs to be inadvertently joined into a single PDU and passed up to the next layer in the protocol stack. The buffer allocation size (BASize) field is not necessarily the length of the PDU (which appears in the trailer); it is supposed to be a hint to the reassembly process as to how much buffer space to allocate for the reassembly. The reason for not including CPI Btag BASize Pad 0 Etag Len 816 0–24 8 816< 64 KB8 User data Figure 3.18 ATM Adaptation Layer 3/4 packet format. 200 3 Packet Switching ATM header Length CRC-10 40 2 4 SEQ MIDType Payload 352 (44 bytes)10 6 10 Figure 3.19 ATM cell format for AAL3/4. Value Name Meaning 10 BOM Beginning of message 00 COM Continuation of message 01 EOM End of message 11 SSM Single-segment message Table 3.4 AAL3/4 Type field. the actual length here is that the sending host might not have known how long the CS-PDU was when it transmitted the header. Before adding the CS-PDU trailer, the user data is padded to one byte less than a multiple of 4 bytes, by adding up to 3 bytes of padding. This padding, plus the 0-filled byte, ensures that the trailer is aligned on a 32-bit boundary, making for more efficient processing. The CS-PDU trailer itself contains the Etag and the real length of the PDU (Len). In addition to the CS-PDU header and trailer, AAL3/4 specifies a header and trailer that are carried in each cell, as depicted in Figure 3.19. Thus, the CS-PDU is actually segmented into 44-byte chunks; an AAL3/4 header and trailer is attached to each one, bringing it up to 48 bytes, which is then carried as the payload of an ATM cell. The first two bits of the AAL3/4 header contain the Type field, which indicates if this is the first cell of a CS-PDU, the last cell of a CS-PDU, a cell in the middle of a CS-PDU, or a single-cell PDU (in which case it is both first and last). The official names for these four conditions are shown in Table 3.4, along with the bit encodings. Next is a 4-bit sequence number (SEQ), which is intended simply to detect cell loss or misordering so that reassembly can be aborted. Clearly, a sequence number this small can miss cell losses if the number of lost cells is large enough. This is followed by a multiplexing identifier (MID), which can be used to multiplex several PDUs onto a single connection. The 6-bit Length field shows the number of bytes of PDU that are contained in the cell; it must equal 44 for BOM and COM cells. Finally, a 10-bit CRC is used to detect errors anywhere in the 48-byte cell payload. 3.3 Cell Switching (ATM) 201 CS-PDU header CS-PDU trailerUser data 44 bytes 44 bytes 44 bytes ≤ 44 bytes ATM header AAL header Cell payload AAL trailer Padding Figure 3.20 Encapsulation and segmentation for AAL3/4. Figure 3.20 shows the entire encapsulation and segmentation process for AAL3/4. At the top, the user data is encapsulated with the CS-PDU header and trailer. The CS-PDU is then segmented into 44-byte payloads, which are encapsulated as ATM cells by adding the AAL3/4 header and trailer as well as the 5-byte ATM header. Note that the last cell is only partially filled whenever the CS-PDU is not an exact multiple of 44 bytes. One thing to note about AAL3/4 is that it exacerbates the fixed per-cell overhead that we discussed above. With 44 bytes of data to 9 bytes of header, the best possible bandwidth utilization would be 83%. Note that the efficiency can be considerably less than that, as illustrated by Figure 3.20, because of the CS-PDU encapsulation and the partial filling of the last cell. ATM Adaptation Layer 5 One thing you may have noticed in the discussion of AAL3/4 is that it seems to take a lot of fields and thus a lot of overhead to perform the conceptually simple function of segmentation and reassembly. This observation was, in fact, made by several people in the early days of ATM, and numerous competing proposals arose for an AAL to support computer communications over ATM. There was a movement, known informally as “Back the Bit,” that argued that if we could just have 1 bit in the ATM header (as opposed to the AAL header) to delineate the end of a frame, then segmentation and reassembly could be accomplished without using any of the 48-byte ATM payload for segmentation/reassembly information. This movement eventually led to the definition of the user signalling bit described above and to the standardization of AAL5. 202 3 Packet Switching CRC-32 < 64 KB 0–47 bytes 16 16 ReservedPad Len 32 Data Figure 3.21 ATM Adaptation Layer 5 packet format. User data 48 bytes 48 bytes 48 bytes ATM header Cell payload Padding CS-PDU trailer Figure 3.22 Encapsulation and segmentation for AAL5. What AAL5 does is replace the 2-bit Type field of AAL3/4 with 1 bit of framing information in the ATM cell header. By setting that 1 bit, we can identify the last cell of a PDU; the next cell is assumed to be the first cell of the next PDU, and subsequent cells are assumed to be COM cells until another cell is received with the user sig- nalling bit set. All the pieces of AAL3/4 that provide protection against lost, corrupt, or misordered cells, including the loss of an EOM cell, are provided by the AAL5 CS-PDU packet format depicted in Figure 3.21. The AAL5 CS-PDU consists simply of the data portion (the PDU handed down by the higher-layer protocol) and an 8-byte trailer. To make sure that the trailer always falls at the tail end of an ATM cell, there may be up to 47 bytes of padding between the data and the trailer. It is necessary to force the trailer to be at the end of a cell, as otherwise there would be no way for the entity performing reassembly of the CS-PDU to find the trailer. The first 2 bytes of the trailer are currently reserved and must be 0. The length field (Len) is the number of bytes carried in the PDU, not including the trailer or any padding before the trailer. Finally, there is a 32-bit CRC. Figure 3.22 shows the encapsulation and segmentation process for AAL5. Just like AAL3/4, the user data is encapsulated to form a CS-PDU (although using only a 3.3 Cell Switching (ATM) 203 trailer in this case). The resulting PDU is then cut up into 48-byte chunks, which are carried directly inside the payload of ATM cells without any further encapsulation. Somewhat surprisingly, AAL5 provides almost the same functionality as AAL3/4 without using an extra 4 bytes out of every cell. For example, the CRC-32 detects lost or misordered cells as well as bit errors in the data. In fact, having a checksum over the entire PDU rather than doing it on a per-cell basis as in AAL3/4 provides stronger protection. For example, it protects against the loss of 16 consecutive cells, an event that would not be picked up by the sequence number checking of AAL3/4. Also, a 32-bit CRC protects against longer burst errors than a 10-bit CRC. The main feature missing from AAL5 is the ability to provide an additional layer of multiplexing onto one virtual circuit using the MID. It is not clear whether this is a significant loss. It is still possible to multiplex traffic from many applications and higher-layer protocols onto a single VC using AAL5 by carrying a demux key of the sort described in Section 1.3.1. It just becomes necessary to do the multiplexing on a packet-by-packet, rather than a cell-by-cell, basis. There are positive and negative aspects to multiplexing traffic from a lot of different applications onto a single VC. For example, if you are being charged for every virtual circuit you set up across a network, then multiplexing traffic from lots of different applications onto one connection might be a plus. However, this approach has the drawback that all applications will have to live with whatever quality of service (e.g., delay and bandwidth guarantees) has been chosen for that one connection, which may mean that some applications are not receiving appropriate service. In general, AAL5 has been wholeheartedly embraced by the computer commu- nications community (at least by that part of the community that has embraced ATM at all). For example, it is the preferred AAL in the IETF for transmitting IP datagrams over ATM. Its more efficient use of bandwidth and simple design are the main features that make it more appealing than AAL3/4. 3.3.3 Virtual Paths As mentioned above, ATM uses a 24-bit identifier for virtual circuits, and these circuits operate almost exactly like the ones described in Section 3.1.2. The one twist is that the 24-bit identifier is split into two parts: an 8-bit virtual path identifier (VPI) and a 16-bit virtual circuit identifier (VCI). This effectively creates a two-level hierarchy of virtual connections. To understand how such a hierarchy might work, consider the following example. (We ignore the fact that in some places there might be a network-network interface with a different-sized VPI; just assume that 8-bit VPIs are used everywhere.) Suppose that a corporation has two sites that connect to a public ATM network, and that at each site the corporation has a network of ATM switches. We could imagine establishing a virtual path between two sites using only the VPI field. Thus, the switches 204 3 Packet Switching Public network Network BNetwork A Figure 3.23 Example of a virtual path. in the public network would use the VPI as the only field on which to make forwarding decisions. From their point of view, this is a virtual circuit network with 8-bit circuit identifiers. The 16-bit VCI is of no interest to these public switches, and they neither use the field for switching nor remap it. Within the corporate sites, however, the full 24-bit space is used for switching. Any traffic that needs to flow between the two sites is routed to a switch that has a connection to the public network, and its top 8 bits (the VPI) are mapped onto the appropriate value to get the data to the other site. This idea is illustrated in Figure 3.23. Note that the virtual path acts like a fat pipe that contains a bundle of virtual circuits, all of which have the same 8 bits in their most significant byte. The advantage of this approach is clear: Although there may be thousands or millions of virtual connections across the public network, the switches in the public network behave as if there is only one connection. This means that there needs to be much less connection-state information stored in the switches, avoiding the need for big, expensive tables of per-VCI information. 3.3.4 Physical Layers for ATM While the layered approach to protocol design might lead you to think that we do not need to worry about what type of point-to-point link ATM runs on top of, this turns out not to be the case. From a simple pragmatic point of view, when you buy an ATM adaptor for a workstation or an ATM switch, it comes with some physical medium over which ATM cells will be sent. Of course, this is also true for other networking protocols such as 802.5 and Ethernet. Like these protocols, ATM can also run over several different physical media and physical-layer protocols. From early in the process of standardizing ATM, it was assumed that ATM would run on top of a SONET physical layer (see Section 2.3.3). Some people even get ATM and SONET confused because they have been so tightly coupled for so long. While it is true that standard ways of carrying ATM cells inside a SONET frame have been defined, and that you can now buy ATM-over-SONET products, the two are entirely 3.3 Cell Switching (ATM) 205 separable. For example, you can lease a SONET link from a phone company and send whatever you want over it, including variable-length packets. Also, you can send ATM cells over many other physical layers instead of SONET, and standards have been (or are being) defined for these encapsulations. A notable early physical layer for ATM was TAXI, the physical layer used in FDDI (Section 2.7). Wireless physical layers for ATM are also being defined. When you send ATM cells over some physical medium, the main issue is how to find the boundaries of the ATM cells; this is exactly the framing problem described in Chapter 2. With SONET, there are two easy ways to find the boundaries. One of the overhead bytes in the SONET frame can be used as a pointer into the SONET payload to the start of an ATM cell. Having found the start of one cell, it is known that the next cell starts 53 bytes further on in the SONET payload, and so on. In theory, you only need to read this pointer once, but in practice, it makes sense to read it every time the SONET overhead goes by so that you can detect errors or resynchronize if needed. The other way to find the boundaries of ATM cells takes advantage of the fact that every cell has a CRC in the fifth byte of the cell. Thus, if you run a CRC calculation over the last 5 bytes received and the answer comes out to indicate no errors, then it is probably true that you have just read an ATM header. If this happens several times in a row at 53-byte intervals, you can be pretty sure you have found the cell boundary. 3.3.5 ATM in the LAN As we mentioned above, ATM grew out of the telephony community, who envisioned it as a way to build large public networks that could transport voice, video, and data traffic. However, it was subsequently embraced by segments of the computer and data communications industries as a technology to be used in LANs—a replacement for Ethernet and 802.5. Its popularity in this realm at a particular point in time can be attributed to two main factors: ■ ATM is a switched technology, whereas Ethernet and 802.5 were originally envisioned as shared-media technologies. ■ ATM was designed to operate on links with speeds of 155 Mbps and above, compared to the original 10 Mbps of Ethernet and 4 or 16 Mbps of token rings. When ATM switches first became available, these were significant advantages over the existing solutions. In particular, switched networks have a big performance advantage over shared-media networks: A single shared-media network has a fixed total bandwidth that must be shared among all hosts, whereas each host gets its own 206 3 Packet Switching dedicated link to the switch in a switched network. Thus the performance of switched networks scales better than that of shared-media networks. However, it should be apparent that the distinction between shared-media and switched networks is not all that clear-cut. A bridge that connects a number of shared- media networks together is also a switch, and it is possible (and quite common) to connect only one host to each segment, giving it dedicated access to that bandwidth. At the same time as ATM switches were appearing on the scene, high-performance Ethernet switches became available. These devices have large numbers of ports and high total throughput. The 100-Mbps Ethernet standard was defined, and so the link speed of Ethernet—which could be achieved over copper—began to approach that of ATM. All this was not enough to kill off ATM in the LAN. One advantage of ATM over Ethernet that remains is the lack of distance limitation for ATM links. Also, higher-speed ATM links (e.g., 622 Mbps) soon became available. This made ATM fairly popular for the high-performance “backbone” of larger LANs. One common configuration was to connect hosts to Ethernet switches, which in turn could be in- terconnected by ATM switches, as depicted in Figure 3.24. High-performance servers might also be connected directly to the ATM switch, as with host H7 in this example. More recently, the technology that has probably overtaken ATM for LAN back- bones and server connections is Gigabit Ethernet. Gigabit Ethernet links use the same framing as lower-speed Ethernets but are usually point-to-point fiber links and can run over relatively long distances (up to several kilometers). And the same basic approach is now scaling up to provide 10-Gbps links. ATM links Ethernet links Ethernet switch ATM switch ATM-attached host E1 H5 H6 H7 H1 E3 H2 H4 H3 E2 Figure 3.24 ATM used as a LAN backbone. 3.3 Cell Switching (ATM) 207 One significant problem with running ATM in a LAN is that it doesn’t look like a “traditional” LAN. Because most LANs (i.e., Ethernets and token rings) are shared- media networks (i.e., every node on the LAN is connected to the same link), it is easy to implement broadcast (sending to everybody) and multicast (sending to a group). Thus, many of the protocols that people depend on in their LANs—for example, the Address Resolution Protocol (ARP) described in Section 4.1.5—depend in turn on the ability of the LAN to support multicast and broadcast. However, because of its connection-oriented and switched nature, ATM behaves rather differently than a shared-media LAN. For example, how can you broadcast to all nodes on an ATM LAN if you don’t know all their addresses and set up VCs to all of them? There are two possible solutions to this problem, and both of them have been explored. One is to redesign the protocols that make assumptions about LANs that are not in fact true of ATM. Thus, for example, there is a new protocol called ATMARP that, unlike traditional ARP, does not depend on broadcast. We discuss this in Section 4.1.5. The alternative is to make ATM behave more like a shared-media LAN—in the sense of supporting multicast and broadcast—without losing the per- formance advantages of a switched network. This approach has been specified by the ATM Forum as “LAN emulation” or LANE (which might be more correctly called “shared-media emulation”). This approach aims to add functionality to ATM LANs so that anything that runs over a shared-media LAN can operate over an ATM LAN. While LANE might now be considered something of a historical curiosity, it does pro- vide an interesting case study in how layering can work in a network. By making the “ATM layer” look more like an Ethernet, higher-layer protocols that worked well over Ethernet continue to work without modification. One aspect of LAN emulation that can be confusing is the variety of different addresses and identifiers that are used. All ATM devices must have an ATM address, which is used when signalling to establish a VC. As noted above, these addresses are different from the standard IEEE 802 MAC addresses used in Ethernets, token rings, and so on. If we want to emulate the behavior of these types of LANs, each device will also need to have a standard (48-bit, globally unique) MAC address. And finally, recall that a virtual circuit identifier is very different from an address. It is the shorthand that is used to get cells along an established connection, but you must first establish a connection, and to do that you need an ATM address. LAN emulation does not actually change the functionality of ATM switches, but adds functionality to the network through the addition of a number of servers. Devices that connect to the ATM network—hosts, bridges, routers—are referred to as LAN emulation clients (LECs). The interactions between LECs and the various servers result in network behavior that, from the point of view of any higher-layer protocol, is indistinguishable from that of an Ethernet or token ring network. 208 3 Packet Switching Host Switch Host Ethernet-like interface Higher-layer protocols (IP, ARP, . . .) Signalling + LANE AAL5 ATM PHY Higher-layer protocols (IP, ARP, . . .) Signalling + LANE AAL5 ATM PHY ATM PHY PHY Figure 3.25 Protocol layers in LAN emulation. Figure 3.25 illustrates the protocol layers in the case where a pair of hosts communi- cate across an ATM network that is emulating a LAN. By “Ethernet-like interface,” we mean that the services offered up to higher layers are like those of an Ethernet: Frames can be delivered to any MAC address on the LAN, frames can be broadcast to all destinations on the LAN, and so on. The servers that are required to build an emulated LAN are ■ the LAN emulation configuration server (LECS) ■ the LAN emulation server (LES) ■ the broadcast and unknown server (BUS) These servers can be physically located in one or more devices, perhaps in one of the hosts or other devices connected to the ATM network. The LECS and LES primarily perform configuration functions, while the BUS has a central role in making data transfer in an ATM network resemble that of a shared-media LAN. The LECS enables a newly attached or rebooted LAN emulation client (e.g., a host) to get some essential information. First, the client must find the LECS, which it may do by using a well-known, predefined VC that is always set up; alternatively, the client must have prior knowledge of the ATM address of the LECS so it can set up a VC to it. Once connected to the LECS, the client provides the LECS with its ATM address, and the LECS responds by telling the client what type of LAN is being emulated (Ethernet or token ring), what the maximum packet size is, and the ATM address of the LES. One LECS might support many separate emulated LANs. 3.3 Cell Switching (ATM) 209 BUSLES ATM network Point-to-point VC Point-to-multipoint VC H1 H2 Figure 3.26 Servers and clients in an emulated LAN. The client now signals for a connection to the LES whose ATM address it just learned. Once connected to the LES, the client registers its MAC and ATM addresses with the LES. Among other things, the LES provides the client with the ATM address of the BUS. The BUS maintains a single point-to-multipoint VC that connects it to all regis- tered clients. It should be apparent that the BUS and this multipoint VC are crucial to LAN emulation: They enable the broadcast capability of traditional LANs to be em- ulated in a virtual circuit environment. Once a LEC has the ATM address of the BUS, it signals for a connection to the BUS. The BUS in turn adds the LEC to the point-to- multipoint VC. At this point, everything is ready for the LEC to participate in data transfer. The arrangement where two hosts have connected to the LES and the BUS, and the BUS has formed the point-to-multipoint VC to both of them, is shown in Figure 3.26. The LECS is not shown. This might seem like a lot of work to get the LEC connected to the BUS, but the separation of functions among servers is helpful from a network management standpoint. For example, a great deal of information can be centralized in a single LECS rather than having to be distributed to many LESs, and the amount of special configuration needed in each host is kept to a bare minimum. It should be clear that the BUS is the place to send any packet that needs to be broadcast to all clients on the LAN. While it could also be used for delivery of unicast packets, this would be inefficient. Delivery of unicast packets operates as follows. Assume that a host has a packet that it wants to deliver to a particular MAC address. In a traditional LAN, the packet could be placed on the wire and would be picked up by the intended recipient. In an emulated LAN, the packet needs to be delivered to the recipient over a virtual circuit. But a newly attached host would only have a VC to the LES and the BUS, not the recipient. To make matters worse, it would not even know the ATM address of the recipient, which is required to set up a VC. Thus, the host performs the following steps: 210 3 Packet Switching ■ It sends the packet to the BUS, which it knows can deliver the packet to the destination using its point-to-multipoint VC. ■ It sends an “address resolution” request to the LES, of the form “What ATM address corresponds to this MAC address?” Since all clients should have registered their MAC and ATM addresses with the LES, the LES should be able to answer the query and provide an ATM address to the client. The client can now signal for a VC to the recipient, which it may use to forward subsequent frames to the destination. The reason for using the BUS to send the first packet is to minimize delay, since it may take some time to get a response from the LES and establish a VC. One detail in this process is that LANs are not supposed to deliver frames out of order, and an emulated LAN should be no different. But if some frames are sent via the BUS and then later frames are sent on a direct connection, misordering may occur. LAN emulation procedures include a “flush” mechanism to ensure that the last packet sent down one path has arrived before another one is sent on a new path, thus ensuring in-order delivery. With the above process, a client would eventually end up with direct VCs to all destinations that it has ever sent data to. This might be an excessive number of VCs, and so a client may use a caching algorithm to dispose of VCs that are no longer carrying traffic. A “cache miss” (i.e., the arrival of a packet that needs to be sent to a destination for which no VC exists) will be handled by sending the packet to the BUS. 3.4 Implementation and Performance So far, we have talked about what a switch must do without discussing how to do it. There is a very simple way to build a switch: Buy a general-purpose workstation and equip it with a number of network interfaces. Such a device, running suitable software, can receive packets on one of its interfaces, perform any of the switching functions described above, and send packets out another of its interfaces. This is, in fact, a popular way to build experimental switches when you want to be able to do things like develop new routing protocols, because it offers extreme flexibility and a familiar programming environment. It is also not too far removed from the architecture of many low-end routers (which, as we will see in the next chapter, have much in common with switches). Figure 3.27 shows a workstation with three network interfaces used as a switch. The figure shows a path that a packet might take from the time it arrives on interface 1 until it is output on interface 2. We have assumed here that the workstation has a mechanism to move data directly from an interface to its main memory without having 3.4 Implementation and Performance 211 I/O bus Interface 1 Interface 2 Interface 3 CPU Main memory Figure 3.27 A workstation used as a packet switch. to be directly copied by the CPU, that is, direct memory access (DMA) as described in Section 2.9. Once the packet is in memory, the CPU examines its header to determine which interface the packet should be sent out on. It then uses DMA to move the packet out to the appropriate interface. Note that Figure 3.27 does not show the packet going to the CPU because the CPU inspects only the header of the packet; it does not have to read every byte of data in the packet. The main problem with using a workstation as a switch is that its performance is limited by the fact that all packets must pass through a single point of contention: In the example shown, each packet crosses the I/O bus twice and is written to and read from main memory once. The upper bound on aggregate throughput of such a device (the total sustainable data rate summed over all inputs) is, thus, either half the main memory bandwidth or half the I/O bus bandwidth, whichever is less. (Usually, it’s the I/O bus bandwidth.) For example, a workstation with a 33-MHz, 32-bit-wide I/O bus can transmit data at a peak rate of a little over 1 Gbps. Since forwarding a packet involves crossing the bus twice, the actual limit is 500 Mbps, which is enough to support five 100-Mbps Ethernet interface cards. In practice, the peak bus bandwidth isn’t sustainable, so it’s more likely such a workstation would support only three or four such interface cards. Moreover, this upper bound also assumes that moving data is the only problem— a fair approximation for long packets but a bad one when packets are short. In the latter case, the cost of processing each packet—parsing its header and deciding which output link to transmit it on—is likely to dominate. Suppose, for example, that a worksta- tion can perform all the necessary processing to switch 500,000 packets each second. 212 3 Packet Switching This is sometimes called the packet per sec- ond (pps) rate. (This number is representa- tive of what is achievable on today’s high- end PCs.) If the average packet is short, say, 64 bytes, this would imply Throughput = pps × (BitsPerPacket) = 500 × 103 × 64 × 8 = 256 × 106 that is, a throughput of 256 Mbps— substantially below the range that users are demanding from their networks today. Bear in mind that this 256 Mbps would be shared by all users connected to the switch, just as the 10 Mbps of an Ethernet is shared among all users connected to the shared medium. Thus, for example, a 10-port switch with this aggregate throughput would only be able to cope with an average data rate of 25.6 Mbps on each port. To address this problem, hardware de- signers have come up with a large array of switch designs that reduce the amount of contention and provide high aggregate throughput. Note that some contention is unavoidable: If every input has data to send to a single output, then they cannot all send it at once. However, if data destined for dif- ferent outputs is arriving at different inputs, a well-designed switch will be able to move data from inputs to outputs in parallel, thus increasing the aggregate throughput. 3.4.1 Ports Most switches look conceptually similar to the one shown in Figure 3.28. They con- sist of a number of input ports and output ports and a fabric. There is usually at least Defining Throughput It turns out to be difficult to de- fine precisely the throughput of a switch. Intuitively, we might think that if a switch has n inputs that each support a link speed of si , then the throughput would just be the sum of all the si . This is actually the best possible throughput that such a switch could provide, but in practice almost no real switch can guarantee that level of perfor- mance. One reason for this is sim- ple to understand. Suppose that, for some period of time, all the traf- fic arriving at the switch needed to be sent to the same output. As long as the bandwidth of that out- put is less than the sum of the input bandwidths, then some of the traf- fic will need to be either buffered or dropped. With this particular traffic pattern, the switch could not provide a sustained through- put higher than the link speed of that one output. However, a switch might be able to handle traffic ar- riving at the full link speed on all inputs if it is distributed across all the outputs evenly; this would be considered optimal. Another factor that affects the performance of switches is the the size of packets arriving on the inputs. For an ATM switch, this is normally not an issue because 3.4 Implementation and Performance 213 all “packets” (cells) are the same length. But for Ethernet switches or IP routers, packets of widely varying sizes are possible. Some of the operations that a switch must perform have a constant overhead per packet, so a switch is likely to perform differently depending on whether all arriving packets are very short, very long, or mixed. For this reason, routers or switches that forward variable-length pack- ets are often characterized by a packet per second (pps) rate as well as a throughput in bits per second. The pps rate is usually measured with minimum-sized packets. The first thing to notice about this discussion is that the through- put of the switch is a function of the traffic to which it is subjected. One of the things that switch designers spend a lot of their time doing is trying to come up with traffic mod- els that approximate the behavior of real data traffic. It turns out that it is extremely difficult to achieve accurate models. A traffic model at- tempts to answer several important questions: (1) When do packets ar- rive? (2) What outputs are they des- tined for? And (3) how big are they? Traffic modeling is a well- established science that has been extremely successful in the world of telephony, enabling telephone com- panies to engineer their networks one control processor in charge of the whole switch that communicates with the ports either directly or, as shown here, via the switch fabric. The ports communicate with the outside world. They may contain fiber optic receivers and lasers, buffers to hold packets that are waiting to be switched or transmitted, and often a significant amount of other circuitry that enables the switch to function. The fabric has a very simple and well-defined job: When presented with a packet, deliver it to the right output port. One of the jobs of the ports, then, is to deal with the complexity of the real world in such a way that the fabric can do its relatively simple job. For example, suppose that this switch is supporting a virtual cir- cuit model of communication. In general, the virtual circuit mapping tables described in Section 3.1.2 are located in the ports. The ports maintain lists of virtual circuit identifiers that are currently in use, with information about what output a packet should be sent out on for each VCI and how the VCI needs to be remapped to ensure uniqueness on the outgoing link. Similarly, the ports of an Ethernet switch store tables that map between Ethernet addresses and output ports (bridge forwarding tables as described in Section 3.2). In general, when a packet is handed from an input port to the fabric, the port has figured out where the packet needs to go, and either the port sets up the fabric accordingly by comm- unicating some control information to it, or it attaches enough information to the packet itself (e.g., an output port number) to al- low the fabric to do its job automatically. Fabrics that switch packets by looking only 214 3 Packet Switching at the information in the packet are referred to as “self-routing,” since they require no external control to route packets. An ex- ample of a self-routing fabric is discussed below. The input port is the first place to look for performance bottlenecks. The in- put port has to receive a steady stream of packets, analyze information in the header of each one to determine which output port (or ports) the packet must be sent to, and pass the packet on to the fabric. The type of header analysis that it performs can range from a simple table lookup on a VCI to complex matching algorithms that examine many fields in the header. This is the type of operation that sometimes becomes a prob- lem when the average packet size is very small. For example, 64-byte packets arriv- ing on a port connected to an OC-48 link have to process packets at a rate of 2.48 × 109 ÷ (64 × 8) = 4.83 × 106 pps In other words, when small packets are arriving as fast as possible on this link (the worst-case scenario that most ports are engineered to handle), the input port has approximately 200 nanoseconds to process each packet. Another key function of ports is buffering. Observe that buffering can hap- pen in either the input or the output port; it can also happen within the fab- ric (sometimes called internal buffering). Simple input buffering has some seri- ous limitations. Consider an input buffer to carry expected loads quite effi- ciently. This is partly because the way people use the phone network does not change that much over time: The frequency with which calls are placed, the amount of time taken for a call, and the tendency of everyone to make calls on Mother’s Day have stayed fairly constant for many years.4 By contrast, the rapid evolution of computer communi- cations, where a new application like Napster can change the traffic patterns almost overnight, has made effective modeling of com- puter networks much more diffi- cult. Nevertheless, there are some excellent books and articles on the subject that we list at the end of the chapter. To give you a sense of the range of throughputs that designers need to be concerned about, a high-end router used in the Internet at the time of writing might support 10 OC-192 links for a throughput of approximately 100 Gbps. A 100- Gbps switch, if called upon to han- dle a steady stream of 64-byte pack- ets, would need a packet per second rate of 100 × 109 ÷ (64 × 8) = 195 × 106 pps 4This statement has recently become less true with the advent of fax machines and modem connections to the Internet. 3.4 Implementation and Performance 215 Switch fabric Control processor Output port Input port Figure 3.28 A 4 × 4 switch. Switch 2 21 Port 1 Port 2 Figure 3.29 Simple illustration of head-of-line blocking. implemented as a FIFO. As packets arrive at the switch, they are placed in the input buffer. The switch then tries to forward the packets at the front of each FIFO to their appropriate output port. However, if the packets at the front of several different input ports are destined for the same output port at the same time, then only one of them can be forwarded;5 the rest must stay in their input buffers. The drawback of this feature is that those packets left at the front of the input buffer prevent other packets further back in the buffer from getting a chance to go to their chosen outputs, even though there may be no contention for those outputs. This phenomenon is called head-of-line blocking. A simple example of head-of-line blocking is given in Figure 3.29, where we see a packet destined for port 1 blocked behind a packet contending for port 2. It can be shown that when traffic is uni- formly distributed among outputs, head-of-line blocking limits the throughput of an 5For a simple input-buffered switch, exactly one packet at a time can be sent to a given output port. It is possible to design switches that can forward more than one packet to the same output at once, at a cost of higher switch complexity, but there is always some upper limit on the number. 216 3 Packet Switching input-buffered switch to 59% of the theoretical maximum (which is the sum of the link bandwidths for the switch). Thus, the majority of switches use either pure out- put buffering or a mixture of internal and output buffering. Those that do rely on input buffers use sophisticated buffer management schemes to avoid head-of-line blocking. Buffers actually perform a more complex task than just holding onto packets that are waiting to be transmitted. Buffers are the main source of delay in a switch, and also the place where packets are most likely to get dropped due to lack of space to store them. The buffers therefore are the main place where the quality of service characteristics of a switch are determined. For example, if a certain packet has been sent along a VC that has a guaranteed delay, it cannot afford to sit in a buffer for very long. This means that the buffers, in general, must be managed using packet scheduling and discard algorithms that meet a wide range of QoS requirements. We talk more about these issues in Chapter 6. 3.4.2 Fabrics While there has been an abundance of impressive research conducted on the design of efficient and scalable fabrics, it is sufficient for our purposes here to understand only the high-level properties of a switch fabric. A switch fabric should be able to move packets from input ports to output ports with minimal delay and in a way that meets the throughput goals of the switch. That usually means that fabrics display some degree of parallelism. A high-performance fabric with n ports can often move one packet from each of its n ports to one of the output ports at the same time. A sample of fabric types includes the following: ■ Shared-bus: This is the type of “fabric” found in a conventional workstation used as a switch, as described above. Because the bus bandwidth determines the throughput of the switch, high-performance switches usually have spe- cially designed busses rather than the standard busses found in PCs. ■ Shared-memory: In a shared-memory switch, packets are written into a mem- ory location by an input port and then read from memory by the output ports. Here it is the memory bandwidth that determines switch throughput, so wide and fast memory is typically used in this sort of design. A shared-memory switch is similar in principle to the shared-bus switch, except it usually uses a specially designed, high-speed memory bus rather than an I/O bus. ■ Crossbar: A crossbar switch is a matrix of pathways that can be configured to connect any input port to any output port. Figure 3.30 shows a 4 × 4 3.4 Implementation and Performance 217 Figure 3.30 A 4 × 4 crossbar switch. crossbar switch. The main problem with crossbars is that, in their simplest form, they require each output port to be able to accept packets from all inputs at once, implying that each port would have a memory bandwidth equal to the total switch throughput. In reality, more complex designs are typically used to address this issue (see, for example, the Knockout switch and McKeown’s virtual output-buffered approach in the “Further Reading” section at the end of the chapter). ■ Self-routing: As noted above, self-routing fabrics rely on some information in the packet header to direct each packet to its correct output. Usually, a special “self-routing header” is appended to the packet by the input port af- ter it has determined which output the packet needs to go to, as illustrated in Figure 3.31; this extra header is removed before the packet leaves the switch. Self-routing fabrics are often built from large numbers of very sim- ple 2 × 2 switching elements interconnected in regular patterns, such as the banyan switching fabric shown in Figure 3.32. For some examples of self- routing fabric designs, see the “Further Reading” section at the end of this chapter. Self-routing fabrics are among the most scalable approaches to fabric design, and there has been a wealth of research on the topic, some of which is listed in the “Further Reading” section. Many self-routing fabrics resemble the one shown in Figure 3.32, 218 3 Packet Switching Switch fabric Output port Input port Original packet header Switch fabric Output port Input port Self-routing header Switch fabric Output port Input port (a) (b) (c) Figure 3.31 A self-routing header is applied to a packet at input to enable the fabric to send the packet to the correct output, where it is removed. (a) Packet arrives at input port. (b) Input port attaches self-routing header to direct packet to correct output. (c) Self-routing header is removed at output port before packet leaves switch. consisting of regularly interconnected 2×2 switching elements. For example, the 2×2 switches in the banyan network perform a simple task: They look at 1 bit in each self- routing header and route packets toward the upper output if it is 0 or toward the lower output if it is 1. Obviously, if two packets arrive at a banyan element at the 3.4 Implementation and Performance 219 001 011 110 111 001 011 110 111 Figure 3.32 Routing packets through a banyan network. The 3-bit numbers represent values in the self-routing headers of four arriving packets. same time and both have the bit set to the same value, then they want to be routed to the same output and a collision will occur. Either preventing or dealing with these collisions is a main challenge for self-routing switch design. The banyan network is a clever arrangement of 2 × 2 switching elements that routes all packets to the correct output without collisions if the packets are presented in ascending order. We can see how this works in an example, as shown in Figure 3.32, where the self-routing header contains the output port number encoded in binary. The switch elements in the first column look at the most significant bit of the output port number and route packets to the top if that bit isa0orthebottom if it is a 1. Switch elements in the second column look at the second bit in the header, and those in the last column look at the least significant bit. You can see from this example that the packets are routed to the correct destination port without collisions. Notice how the top outputs from the first column of switches all lead to the top half of the network, thus getting packets with port numbers 0–3 into the right half of the network. The next column gets packets to the right quarter of the network, and the final column gets them to the right output port. The clever part is the way switches are arranged to avoid collisions. Part of the arrangement includes the “perfect shuffle” wiring pattern at the start of the network. To build a complete switch fabric around a banyan network would require additional components to sort packets before they are presented to the banyan. The Batcher-banyan switch design is a notable example of such an approach. The Batcher network, which is also built from a regular interconnection of 2×2 switching elements, 220 3 Packet Switching sorts packets into descending order. On leaving the Batcher network, the packets are then ready to be directed to the correct output, with no risk of collisions, by the banyan network. One of the interesting things about switch design is the wide range of different types of switches that can be built using the same basic technology. For example, the Ethernet switches and ATM switches discussed in this chapter, as well as Internet routers discussed in the next chapter, are all built using designs such as those outlined in this section. 3.5 Summary This chapter has started to look at some of the issues involved in building large scalable networks by using switches, rather than just links, to interconnect hosts. There are several different ways to decide how to switch packets; the two main ones are the datagram (connectionless) model and the virtual circuit (connection-oriented) model. An important application of switching is the interconnection of shared-media LANs. LAN switches, or bridges, use techniques such as source address learning to improve forwarding efficiency and spanning tree algorithms to avoid looping. These switches are extensively used in data centers, campuses, and corporate networks. The most widespread uses of virtual circuit switching are in Frame Relay and ATM switches. ATM introduces some particular challenges through the use of cells, short fixed-length packets. The availability of relatively high-throughput ATM switches has contributed to the acceptance of the technology, although it has certainly not swept all other technologies aside as some predicted. One of the main uses of ATM today is to interconnect widely separated sites in corporate networks. Independent of the specifics of the switching technology, switches need to forward packets from inputs to outputs at a high rate, and in some circumstances, switches need to grow to a large size to accommodate hundreds or thousands of ports. Building switches that both scale and offer high performance at acceptable cost is complicated by the problem of contention, and as a consequence, switches often employ special- purpose hardware rather than being built from general-purpose workstations. In addition to the issues of contention discussed here, we observe that the related problem of congestion has come up throughout this chapter. We will postpone our discussion of congestion control until Chapter 6, after we have seen more of the network architecture. We do this because it is impossible to fully appreciate congestion (both the problem and how it to address it) without understanding both what happens inside the network (the topic of this and the next chapter) and what happens at the edges of the network (the topic of Chapter 5). Further Reading 221 ATM was originally envisioned by many of its proponents as the foun- dation for the “Broadband Integrated Services Digital Network,” and it was predicted in some quarters that ATM would displace all other networking technologies. Hosts would acquire OPEN ISSUE The Future of ATM ATM adaptors instead of Ethernet ports, enabling “ATM to the desktop.” Phone companies everywhere would deploy ATM, and as the technology that supports all media types—voice, video, and data—it would remove the need for any other type of network. It is now apparent that this scenario is unlikely to play out. The success of Eth- ernet switches in particular has killed off the ATM-to-the-desktop movement. Gigabit Ethernet and 10-Gigabit Ethernet technologies have successfully addressed the need for high-speed connections to servers where ATM might once have been used. In fact, we now hear ATM referred to as a “legacy protocol,” a term that was much used in the heyday of ATM to refer to older protocols. Another factor that has limited the acceptance of ATM has been the success of the Internet. It is now a fact of life that consumers are willing to pay for Internet access, and that means selling a service that delivers IP packets. While ATM can be used to help deliver that service (as it is in many DSL networks, for example), simply selling ATM connections to consumers does not meet their data networking needs. The notable exception is corporate customers looking to interconnect many sites, where an ATM VC may be just the right thing to economically replace a leased line. In fact, this is the primary niche for ATM today—it is used (in conjunction with Frame Relay) to provide wide area virtual circuit services to corporate networking customers. The future of ATM therefore seems to hinge on the future of wide area virtual circuit-based services. These services are unlikely to go away any time soon, but ATM’s role is to some extent being challenged by newer technologies, such as encrypted IP tunnels and Multiprotocol Label Switching (MPLS). Both of these technologies will be described in the next chapter. FURTHER READING The seminal paper on bridges, in particular the spanning tree algorithm, is the article by Perlman below. There is a wealth of survey papers on ATM; the article by Turner, an ATM pioneer, is one of the earliest to propose the use of a cell-based network for integrated services. The third paper describes the Sunshine switch and is especially 222 3 Packet Switching interesting because it provides insights into the important role of traffic analysis in switch design. In particular, the Sunshine designers were among the first to realize that cells were unlikely to arrive at a switch in a totally uncorrelated way and thus were able to factor these correlations into their design. Finally, McKeown’s paper describes an approach to switch design that uses cells internally but has been used commercially as the basis for high-performance routers forwarding variable-length packets. ■ Perlman, R. An algorithm for distributed computation of spanning trees in an extended LAN. Proceedings of the Ninth Data Communications Symposium, pages 44–53, September 1985. ■ Turner, J. S. Design of an integrated services packet network. Proceedings of the Ninth Data Communications Symposium, pages 124–133, September 1985. ■ Giacopelli, J. N., et al. Sunshine: A high-performance self-routing broadband packet-switched architecture. IEEE Journal of Selected Areas in Communi- cations (JSAC) 9(8):1289–1298, October 1991. ■ McKeown, N. The iSLIP scheduling algorithm for input-queued switches. IEEE Transactions on Networking 7(2):188–201, April 1999. A good general overview of bridges can be found in another work by Perlman [Per00]. For a detailed description of many aspects of ATM, with a focus on building real networks, we recommend the book by Ginsburg [Gin99]. Also, as one of the key ATM standards-setting bodies, the ATM Forum produces new specifications for ATM; the User-Network Interface (UNI) specification, version 4.1, is the most recent at the time of this writing. (See the live reference below.) There have been literally thousands of papers published on switch architectures. One early paper that explains Batcher networks well is, not surprisingly, one by Batcher himself [Bat68]. Sorting networks are explained by Drysdale and Young [DY75], and the knockout switch, an interesting form of crossbar switch is described by Yeh et al. [YHA87]. A survey of ATM switch architectures appears in Partridge [Par94], and a good overview of the performance of different switching fabrics can be found in Robertazzi [Rob93]. An example of the design of a switch based on variable-length packets can be found in Gopal and Guerin [GG94]. Optical networking is a rich field in its own right, with its own journals, con- ferences, and so on. We recommend Ramaswami and Sivarajan [RS01] as a good introductory text in that field. An excellent text to read if you want to learn about the mathematical analysis of network performance is by Kleinrock [Kle75], one of the pioneers of the ARPANET. Exercises 223 Many papers have been published on the applications of queuing theory to packet switching. We recommend the article by Paxson and Floyd [PF94] as a significant contribution focused on the Internet, and one by Leland et al. [LTWW94], a paper that introduces the important concept of “long-range dependence” and shows the inadequacy of many traditional approaches to traffic modeling. Finally, we recommend the following live reference: ■ http://www.atmforum.com: current activities of the ATM Forum EXERCISES 1 Using the example network given in Figure 3.33, give the virtual circuit tables for all the switches after each of the following connections is established. Assume that the sequence of connections is cumulative; that is, the first connection is still up when the second connection is established, and so on. Also assume that the VCI assignment always picks the lowest unused VCI on each link, starting with 0. (a) Host A connects to host B. (b) Host C connects to host G. (c) Host E connects to host I. (d) Host D connects to host B. (e) Host F connects to host J. (f) Host H connects to host A. 2 Using the example network given in Figure 3.33, give the virtual circuit tables for all the switches after each of the following connections is established. Assume that the sequence of connections is cumulative; that is, the first connection is still up when the second connection is established, and so on. Also assume that the VCI assignment always picks the lowest unused VCI on each link, starting with 0. (a) Host D connects to host H. (b) Host B connects to host G. (c) Host F connects to host A. (d) Host H connects to host C. (e) Host I connects to host E. (f) Host H connects to host J. 224 3 Packet Switching 0 13 2 0 13 2 0 13 2 Switch 3 Switch 2 Switch 4 Host A Host BHost J Switch 1 Host C Host D Host E Host I 2 13 0 Host H Host F Host G Figure 3.33 Example network for Exercises 1 and 2. 2 3 6 2 8 1 D A F E B C Figure 3.34 Network for Exercise 3. 3 For the network given in Figure 3.34, give the datagram forwarding table for each node. The links are labelled with relative costs; your tables should forward each packet via the lowest-cost path to its destination. 4 Give forwarding tables for switches S1–S4 in Figure 3.35. Each switch should have a “default” routing entry, chosen to forward packets with unrecognized Exercises 225 S1 13 2 A B S2 13 2 S3 13 2 C S4 1 2 DOUT Figure 3.35 Diagram for Exercise 4. S1 13 2 A E B S2 13 3 2 C S3 1 2 D Figure 3.36 Diagram for Exercise 5. Switch S1 Port VCI Port VCI 1231 1123 2132 Switch S2 Port VCI Port VCI 1133 1232 Switch S3 Port VCI Port VCI 1321 1231 Table 3.5 VCI tables for switches in Figure 3.36. destination addresses toward OUT. Any specific destination table entries dupli- cated by the default entry should then be eliminated. 5 Consider the virtual circuit switches in Figure 3.36. Table 3.5 lists, for each switch, what port, VCI (or VCI, interface) pairs are connected to what other. Connec- tions are bidirectional. List all endpoint-to-endpoint connections. 6 In the source routing example of Section 3.1.3, the address received by B is not reversible and doesn’t help B know how to reach A. Propose a modification to the delivery mechanism that does allow for reversibility. Your mechanism should not require giving all switches globally unique names. 226 3 Packet Switching 7 Propose a mechanism that virtual circuit switches might use so that if one switch loses all its state regarding connections, then a sender of packets along a path through that switch is informed of the failure. 8 Propose a mechanism that might be used by datagram switches so that if one switch loses all or part of its forwarding table, affected senders are informed of the failure. 9 The virtual circuit mechanism described in Section 3.1.2 assumes that each link is point-to-point. Extend the forwarding algorithm to work in the case that links are shared-media connections, for example, Ethernet. 10 Suppose, in Figure 3.4, that a new link has been added, connecting switch 3 port 1 (where G is now) and switch 1 port 0 (where D is now); neither switch is “informed” of this link. Furthermore, switch 3 mistakenly thinks that host B is reached via port 1. (a) What happens if host A attempts to send to host B, using datagram forwarding? (b) What happens if host A attempts to connect to host B, using the virtual circuit setup mechanism discussed in the text? 11 Give an example of a working virtual circuit whose path traverses some link twice. Packets sent along this path should not, however, circulate indefinitely. 12 In Section 3.1.2, each switch chose the VCI value for the incoming link. Show that it is also possible for each switch to choose the VCI value for the outbound link, and that the same VCI values will be chosen by each approach. If each switch chooses the outbound VCI, is it still necessary to wait one RTT before data is sent? 13 Given the extended LAN shown in Figure 3.37, indicate which ports are not selected by the spanning tree algorithm. 14 Given the extended LAN shown in Figure 3.37, assume that bridge B1 suffers catastrophic failure. Indicate which ports are not selected by the spanning tree algorithm after the recovery process and a new tree has been formed. 15 Consider the arrangement of learning bridges shown in Figure 3.38. Assuming all are initially empty, give the forwarding tables for each of the bridges B1–B4 after the following transmissions: Exercises 227 A C E D B F H J G I B7 B1 B2 B5 B6 B3 B4 Figure 3.37 Network for Exercises 13 and 14. B3 C A B1 B2 B4 D Figure 3.38 Network for Exercises 15 and 16. ■ A sends to C. ■ C sends to A. ■ D sends to C. Identify ports with the unique neighbor reached directly from that port; that is, the ports for B1 are to be labelled “A” and “B2.” 16 As in the previous problem, consider the arrangement of learning bridges shown in Figure 3.38. Assuming all are initially empty, give the forwarding tables for each of the bridges B1–B4 after the following transmissions: 228 3 Packet Switching Z XB1B2B3 YW Figure 3.39 Diagram for Exercise 17. B1 B3B2 Figure 3.40 Extended LAN for Exercise 18. ■ D sends to C. ■ C sends to D. ■ A sends to C. 17 Consider hosts X, Y, Z, W and learning bridges B1, B2, B3, with initially empty forwarding tables, as in Figure 3.39. (a) Suppose X sends to Z. Which bridges learn where X is? Does Y’s network interface see this packet? (b) Suppose Z now sends to X. Which bridges learn where Z is? Does Y’s network interface see this packet? (c) Suppose Y now sends to X. Which bridges learn where Y is? Does Z’s network interface see this packet? (d) Finally, suppose Z sends to Y. Which bridges learn where Z is? Does W’s network interface see this packet? 18 Give the spanning tree generated for the extended LAN shown in Figure 3.40, and discuss how any ties are resolved. Exercises 229 B1 L M B2 Figure 3.41 Loop for Exercises 19 and 20. 19 Suppose two learning bridges B1 and B2 form a loop as shown in Figure 3.41, and do not implement the spanning tree algorithm. Each bridge maintains a single table of address, interface pairs. (a) What will happen if M sends to L? (b) Suppose a short while later L replies to M. Give a sequence of events that leads to one packet from M and one packet from L circling the loop in opposite directions. 20 Suppose that M in Figure 3.41 sends to itself (this normally would never happen). State what would happen, assuming (a) the bridges’ learning algorithm is to install (or update) the new sourceaddress, interface entry before searching the table for the destination address (b) the new source address was installed after destination address lookup 21 Consider the extended LAN of Figure 3.12. What happens in the spanning tree algorithm if bridge B1 does not participate and (a) simply forwards all spanning tree algorithm messages? (b) drops all spanning tree messages? 22 Suppose some repeaters (hubs), rather than bridges, are connected into a loop. (a) What will happen when somebody transmits? (b) Why would the spanning tree mechanism be difficult or impossible to imple- ment for repeaters? 230 3 Packet Switching (c) Propose a mechanism by which repeaters might detect loops and shut down some ports to break the loop. Your solution is not required to work 100% of the time. 23 Suppose a bridge has two of its ports on the same network. How might the bridge detect and correct this? 24 What percentage of an ATM link’s total bandwidth is consumed by the ATM cell headers? What percentage of the total bandwidth is consumed by all nonpayload bits in AAL3/4 and AAL5, when the user data is 512 bytes long? 25 Explain why AAL3/4 will not detect the loss of 16 consecutive cells of a single PDU. 26 The IP datagram for a TCP ACK message is 40 bytes long: It contains 20 bytes of TCP header and 20 bytes of IP header. Assume that this ACK is traversing an ATM network that uses AAL5 to encapsulate IP packets. How many ATM packets will it take to carry the ACK? What if AAL3/4 is used instead? 27 The CS-PDU for AAL5 contains up to 47 bytes of padding, while the AAL3/4 CS- PDU only contains up to 3 bytes of padding. Explain why the effective bandwidth of AAL5 is always the same as, or higher than, that of AAL3/4, given a PDU of a particular size. 28 How reliable does an ATM connection have to be in order to maintain a loss rate of less than one per million for a higher-level PDU of size 20 cells? Assume AAL5. 29 Assuming the 20-cell AAL5 packet from the previous problem, suppose a final cell is tacked on the end of the PDU, and that this cell is the XOR of all the previous cells in the PDU. This allows recovery from any one lost cell. What cell loss rate now would yield a net one-per-million loss rate for 20-data-cell PDUs? 30 Recall that AAL3/4 has a CRC-10 checksum at the end of each cell, while AAL5 has a single CRC-32 checksum at the end of the PDU. If a PDU is carried in 12 AAL3/4 cells, then AAL3/4 devotes nearly four times as many bits to error detection as AAL5. (a) Suppose errors are known to come in bursts, where each burst is small enough to be confined to a single cell. Find the probability that AAL3/4 fails to detect Exercises 231 an error, given that it is known that exactly two cells are affected. Do the same for three cells. Under these conditions, is AAL3/4 more or less reliable than AAL5? Assume that an N-bit CRC fails to detect an error with probability 1/2N (which is strictly true only when all errors are equally likely). (b) Can you think of any error distribution in which AAL3/4 would be more likely than AAL5 to detect an error? Do you think such circumstances are likely? 31 Cell switching methods essentially always use virtual circuit routing rather than datagram routing. Give a specific argument why this is so. 32 Suppose a workstation has an I/O bus speed of 800 Mbps and memory bandwidth of 2 Gbps. Assuming DMA in and out of main memory, how many interfaces to 45-Mbps T3 links could a switch based on this workstation handle? 33 Suppose a workstation has an I/O bus speed of 1 Gbps and memory bandwidth of 2 Gbps. Assuming DMA in and out of main memory, how many interfaces to 45-Mbps T3 links could a switch based on this workstation handle? 34 Suppose a switch can forward packets at a rate of 100,000 per second, regardless (within limits) of size. Assuming the workstation parameters described in the pre- vious problem, at what packet size would the bus bandwidth become the limiting factor? 35 Suppose that a switch is designed to have both input and output FIFO buffering. As packets arrive on an input port they are inserted at the tail of the FIFO. The switch then tries to forward the packets at the head of each FIFO to the tail of the appropriate output FIFO. (a) Explain under what circumstances such a switch can lose a packet destined for an output port whose FIFO is empty. (b) What is this behavior called? (c) Assuming the FIFO buffering memory can be redistributed freely, suggest a reshuffling of the buffers that avoids the above problem, and explain why it does so. 36 A stage of an n × n banyan network consists of (n/2) 2 × 2 switching elements. The first stage directs packets to the correct half of the network, the next stage to the correct quarter, and so on, until the packet is routed to the correct output. 232 3 Packet Switching Derive an expression for the number of 2 × 2 switching elements needed to make an n × n banyan network. Verify your answer for n = 8. 37 Describe how a Batcher network works. (See the “Further Reading” section.) Explain how a Batcher network can be used in combination with a banyan network to build a switching fabric. 38 An Ethernet switch is simply a bridge that has the ability to forward some num- ber of packets in parallel, assuming the input and output ports are all distinct. Suppose two such N-port switches, for a large value of N, are each able to for- ward individually up to three packets in parallel. They are then connected to one another in series by joining a pair of ports, one from each switch; the joining link is the bottleneck as it can, of course, carry only one packet at a time. (a) Suppose we choose two connections through this combined switch at random. What is the probability that both connections can be forwarded in parallel? Hint: This is the probability that at most one of the connections crosses the link. (b) What if three connections are chosen at random? 39 Suppose a 10-Mbps Ethernet hub (repeater) is replaced by a 10-Mbps switch, in an environment where all traffic is between a single server and N “clients.” Because all traffic must still traverse the server-switch link, nominally there is no improvement in bandwidth. (a) Would you expect any improvement in bandwidth? If so, why? (b) What would your answer be if the original hub were token ring rather than Ethernet? (c) What other advantages and drawbacks might a switch offer versus a hub? This Page Intentionally Left Blank Internetworking Every seeming equality conceals a hierarchy. —Mason Cooley We have now seen how to build a single network using point-to-point links, shared media, and switches. The problem is that lots of people have built networks with these various technologies and they all want to be able to communicate with each other, not just with the other users of a single network. This chapter is about the problem of interconnecting different networks. PROBLEM There Is More Than One Network There are two important prob- lems that must be addressed when connecting networks: heterogeneity and scale. Simply stated, the problem of heterogeneity is that users on one type of network want to be able to communicate with users on other type of networks. To further complicate matters, establishing connectivity between hosts on two different networks may require traversing several other networks in between, each of which may be of yet another type. These different networks may be Ethernets, token rings, point-to-point links, or switched networks of various kinds, and each of them is likely to have its own addressing scheme, media access protocols, service model, and so on. The challenge of heterogeneity is to provide a useful and fairly predictable host-to-host service over this hodgepodge of different networks. To understand the problem of scaling, it is worth considering the growth of the Internet, which has roughly doubled in size each year for 20 years. This sort of growth forces us to face a number of challenges. One of these is routing: How can you find an efficient path through a network with millions, or perhaps billions, of nodes? Closely related to this is the problem of addressing, the task of providing suitable identifiers for all those nodes. This chapter looks at a series of approaches to interconnecting networks and the problems that must be solved. In doing so, we trace the evolution of the TCP/IP 4Internet in an effort to understand the problems of het- erogeneity and scale in detail, along with the general tech- niques that can be applied to them. The first section introduces the Internet Protocol (IP) and shows how it can be used to build a scalable, heteroge- neous internetwork. This section includes a discussion of the Internet’s service model, which is the key to its ability to handle heterogeneity. It also describes how the Internet’s hierarchical addressing scheme has helped the Internet to scale to a modestly large size. A central aspect of building large heterogeneous in- ternetworks is the problem of finding efficient, loop-free paths through the constituent networks. The second sec- tion introduces the principles of routing and explores the scaling issues of routing protocols, using some of the Inter- net’s routing protocols as examples. The third section discusses several of the problems (growing pains) that the Internet has experienced over the past several years and introduces a variety of techniques that have been employed to address these problems. The experience gained from using these techniques has led to the design of a new version of IP, which is IP version 6 (IPv6). Throughout all these discussions, we see the impor- tance of hierarchy in building scalable networks. The chapter concludes by considering a pair of sig- nificant enhancements to the Internet’s capabilities. The first, multicast, is an enhancement of the basic service model. We show how multicast—the ability to deliver packets efficiently to a set of receivers—can be incorpo- rated into an internet, and we describe several of the rout- ing protocols that have been developed to support mul- ticast. The second enhancement, MPLS (Multiprotocol Label Switching), modifies the forwarding mechanism of IP networks. This modification has enabled some changes in the way IP routing is performed and in the services of- fered by IP networks. 236 4 Internetworking 4.1 Simple Internetworking (IP) In the previous chapter, we saw that it was possible to build reasonably large LANs using bridges and LAN switches, but that such approaches were limited in their abil- ity to scale and to handle heterogeneity. In this chapter, we explore some ways to go beyond the limitations of bridged networks, enabling us to build large, highly het- erogeneous networks with reasonably efficient routing. We refer to such networks as internetworks. In the following sections, we make a steady progression toward larger and larger internetworks. We start with the basic functionality of the currently deployed version of the Internet Protocol (IP), and then we examine various techniques that have been developed to extend the scalability of the Internet in Section 4.3. This discussion culminates with a description of IP version 6 (IPv6), also known as the “next-generation” IP. Before delving into the details of an internetworking protocol, however, let’s consider more carefully what the word “internetwork” means. 4.1.1 What Is an Internetwork? We use the term “internetwork,” or sometimes just “internet” with a lowercase i, to refer to an arbitrary collection of networks interconnected to provide some sort of host- to-host packet delivery service. For example, a corporation with many sites might con- struct a private internetwork by interconnecting the LANs at their different sites with point-to-point links leased from the phone company. When we are talking about the widely used, global internetwork to which a large percentage of networks are now con- nected, we call it the “Internet” with a capital I. In keeping with the first-principles ap- proach of this book, we mainly want you to learn about the principles of “lowercase i” internetworking, but we illustrate these ideas with real-world examples from the “big I” Internet. Another piece of terminology that can be confusing is the difference between networks, subnetworks, and internetworks. We are going to avoid subnetworks (or subnets) altogether until Section 4.3. For now, we use network to mean either a directly connected or a switched network of the kind that was discussed in the last two chapters. Such a network uses one technology, such as 802.5, Ethernet, or ATM. An internetwork is an interconnected collection of such networks. Sometimes, to avoid ambiguity, we refer to the underlying networks that we are interconnecting as physical networks. An internet is a logical network built out of a collection of physical networks. In this context, a collection of Ethernets connected by bridges or switches would still be viewed as a single network. Figure 4.1 shows an example internetwork. An internetwork is often referred to as a “network of networks” because it is made up of lots of smaller networks. In this figure, we see Ethernets, an FDDI ring, and a point-to-point link. Each of these 4.1 Simple Internetworking (IP) 237 R2 R1 H4 H5 H3H2H1 Network 2 (Ethernet) Network 1 (Ethernet) H6 Network 4 (point-to-point) H7 R3 H8 Network 3 (FDDI) Figure 4.1 A simple internetwork. Hn = host; Rn = router. is a single-technology network. The nodes that interconnect the networks are called routers. They are also sometimes called gateways, but since this term has several other connotations, we restrict our usage to router. The Internet Protocol is the key tool used today to build scalable, heterogeneous internetworks. It was originally known as the Kahn-Cerf protocol after its inventors. One way to think of IP is that it runs on all the nodes (both hosts and routers) in a collection of networks and defines the infrastructure that allows these nodes and networks to function as a single logical internetwork. For example, Figure 4.2 shows how hosts H1 and H8 are logically connected by the internet in Figure 4.1, including the protocol graph running on each node. Note that higher-level protocols, such as TCP and UDP, typically run on top of IP on the hosts. Most of the rest of this chapter is about various aspects of IP. While it is certainly possible to build an internetwork that does not use IP—for example, Novell created an internetworking protocol called IPX, which was in turn based on the XNS internet designed by Xerox—IP is the most interesting case to study simply because of the size of the Internet. Said another way, it is only the IP Internet that has really faced the issue of scale. Thus it provides the best case study of a scalable internetworking protocol. 238 4 Internetworking R1 R2 R3 H1 H8 ETH FDDI IP ETH TCP FDDI PPP PPP ETH IP ETH TCP IP IP IP Figure 4.2 A simple internetwork, showing the protocol layers used to connect H1 to H8 in Figure 4.1. ETH is the protocol that runs over Ethernet. 4.1.2 Service Model A good place to start when you build an internetwork is to define its service model, that is, the host-to-host services you want to provide. The main concern in defining a service model for an internetwork is that we can provide a host-to-host service only if this service can somehow be provided over each of the underlying physical networks. For example, it would be no good deciding that our internetwork service model was going to provide guaranteed delivery of every packet in 1 ms or less if there were underlying network technologies that could arbitrarily delay packets. The philosophy used in defining the IP service model, therefore, was to make it undemanding enough that just about any network technology that might turn up in an internetwork would be able to provide the necessary service. The IP service model can be thought of as having two parts: an addressing scheme, which provides a way to identify all hosts in the internetwork, and a datagram (con- nectionless) model of data delivery. This service model is sometimes called best effort because, although IP makes every effort to deliver datagrams, it makes no guarantees. We postpone a discussion of the addressing scheme for now and look first at the data delivery model. Datagram Delivery The IP datagram is fundamental to the Internet Protocol. Recall from Section 3.1.1 that a datagram is a type of packet that happens to be sent in a connectionless manner over a network. Every datagram carries enough information to let the network forward the packet to its correct destination; there is no need for any advance setup mechanism to tell the network what to do when the packet arrives. You just send it, and the network makes its best effort to get it to the desired destination. The “best-effort” part means 4.1 Simple Internetworking (IP) 239 that if something goes wrong and the packet gets lost, corrupted, misdelivered, or in any way fails to reach its intended destination, the network does nothing—it made its best effort, and that is all it has to do. It does not make any attempt to recover from the failure. This is sometimes called an unreliable service. Best-effort, connectionless service is about the simplest service you could ask for from an internetwork, and this is a great strength. For example, if you provide best- effort service over a network that provides a reliable service, then that’s fine—you end up with a best-effort service that just happens to always deliver the packets. If, on the other hand, you had a reliable service model over an unreliable network, you would have to put lots of extra functionality into the routers to make up for the deficiencies of the underlying network. Keeping the routers as simple as possible was one of the original design goals of IP. The ability of IP to “run over anything” is frequently cited as one of its most important characteristics. It is noteworthy that many of the technologies over which IP runs today did not exist when IP was invented. So far, no networking technology has been invented that has proven too bizarre for IP; it has even been claimed that IP can run over a network that transports messages using carrier pigeons. Best-effort delivery does not just mean that packets can get lost. Sometimes they can get delivered out of order, and sometimes the same packet can get delivered more than once. The higher-level protocols or applications that run above IP need to be aware of all these possible failure modes. Packet Format Clearly, a key part of the IP service model is the type of packets that can be carried. The IP datagram, like most packets, consists of a header followed by a number of bytes of data. The format of the header is shown in Figure 4.3. Note that we have adopted a different style of representing packets than the one we used in previous chapters. This is because packet formats at the internetworking layer and above, where we will be focusing our attention for the next few chapters, are almost invariably designed to align on 32-bit boundaries to simplify the task of processing them in software. Thus, the common way of representing them (used in Internet Requests for Comments, for example) is to draw them as a succession of 32-bit words. The top word is the one transmitted first, and the leftmost byte of each word is the one transmitted first. In this representation, you can easily recognize fields that are a multiple of 8 bits long. On the odd occasion when fields are not an even multiple of 8 bits, you can determine the field lengths by looking at the bit positions marked at the top of the packet. Looking at each field in the IP header, we see that the “simple” model of best- effort datagram delivery still has some subtle features. The Version field specifies the 240 4 Internetworking Version HLen TOS Length Ident Flags Offset TTL Protocol Checksum SourceAddr DestinationAddr Options (variable) Pad (variable) 048 1619 31 Data Figure 4.3 IPv4 packet header. version of IP. The current version of IP is 4, and it is sometimes called IPv4.1 Observe that putting this field right at the start of the datagram makes it easy for everything else in the packet format to be redefined in subsequent versions; the header processing software starts off by looking at the version and then branches off to process the rest of the packet according to the appropriate format. The next field, HLen, specifies the length of the header in 32-bit words. When there are no options, which is most of the time, the header is 5 words (20 bytes) long. The 8-bit TOS (type of service) field has had a number of different definitions over the years, but its basic function is to allow packets to be treated differently based on application needs. For example, the TOS value might determine whether or not a packet should be placed in a special queue that receives low delay. We discuss the use of this field (and a new name for it) in more detail in Section 6.5.3. The next 16 bits of the header contain the Length of the datagram, including the header. Unlike the HLen field, the Length field counts bytes rather than words. Thus, the maximum size of an IP datagram is 65,535 bytes. The physical network over which IP is running, however, may not support such long packets. For this reason, IP supports a fragmentation and reassembly process. The second word of the header contains information about fragmentation, and the details of its use are presented under “Fragmentation and Reassembly” below. 1The next major version of IP, which is discussed later in this chapter, has a new version number 6 and is known as IPv6. The version number 5 was used for an experimental protocol called ST-II that was not widely used. 4.1 Simple Internetworking (IP) 241 Moving on to the third word of the header, the next byte is the TTL (time to live) field. Its name reflects its historical meaning rather than the way it is commonly used today. The intent of the field is to catch packets that have been going around in routing loops and discard them, rather than let them consume resources indefinitely. Originally, TTL was set to a specific number of seconds that the packet would be allowed to live, and routers along the path would decrement this field until it reached 0. However, since it was rare for a packet to sit for as long as 1 second in a router, and routers did not all have access to a common clock, most routers just decremented the TTL by1as they forwarded the packet. Thus, it became more of a hop count than a timer, which is still a perfectly good way to catch packets that are stuck in routing loops. One subtlety is in the initial setting of this field by the sending host: Set it too high and packets could circulate rather a lot before getting dropped; set it too low and they may not reach their destination. The value 64 is the current default. The Protocol field is simply a demultiplexing key that identifies the higher-level protocol to which this IP packet should be passed. There are values defined for TCP (6), UDP (17), and many other protocols that may sit above IP in the protocol graph. The Checksum is calculated by considering the entire IP header as a sequence of 16-bit words, adding them up using ones complement arithmetic, and taking the ones complement of the result. This is the IP checksum algorithm described in Section 2.4. Thus, if any bit in the header is corrupted in transit, the checksum will not contain the correct value upon receipt of the packet. Since a corrupted header may contain an error in the destination address—and, as a result, may have been misdelivered—it makes sense to discard any packet that fails the checksum. It should be noted that this type of checksum does not have the same strong error detection properties as a CRC, but it is much easier to calculate in software. The last two required fields in the header are the SourceAddr and the DestinationAddr for the packet. The latter is the key to datagram delivery: Every packet contains a full address for its intended destination so that forwarding decisions can be made at each router. The source address is required to allow recipients to decide if they want to accept the packet and to enable them to reply. IP addresses are discussed in Section 4.1.3—for now, the important thing to know is that IP defines its own global address space, independent of whatever physical networks it runs over. As we will see, this is one of the keys to supporting heterogeneity. Finally, there may be a number of options at the end of the header. The presence or absence of options may be determined by examining the header length (HLen) field. While options are used fairly rarely, a complete IP implementation must handle them all. 242 4 Internetworking Fragmentation and Reassembly One of the problems of providing a uniform host-to-host service model over a het- erogeneous collection of networks is that each network technology tends to have its own idea of how large a packet can be. For example, an Ethernet can accept packets up to 1500 bytes long, while FDDI packets may be 4500 bytes long. This leaves two choices for the IP service model: make sure that all IP datagrams are small enough to fit inside one packet on any network technology, or provide a means by which packets can be fragmented and reassembled when they are too big to go over a given network technology. The latter turns out to be a good choice, especially when you consider the fact that new network technologies are always turning up, and IP needs to run over all of them; this would make it hard to pick a suitably small bound on datagram size. This also means that a host will not send needlessly small packets, which wastes bandwidth and consumes processing resources by requiring more headers per byte of data sent. For example, two hosts connected to FDDI networks that are interconnected by a point-to-point link would not need to send packets small enough to fit on an Ethernet. The central idea here is that every network type has a maximum transmission unit (MTU), which is the largest IP datagram that it can carry in a frame. Note that this value is smaller than the largest packet size on that network because the IP datagram needs to fit in the payload of the link-layer frame. Also, note that in ATM networks, the “frame” is the CS-PDU, not the ATM cell; the fact that CS-PDUs get segmented into cells is not visible to IP. When a host sends an IP datagram, therefore, it can choose any size that it wants. A reasonable choice is the MTU of the network to which the host is directly attached. Then fragmentation will only be necessary if the path to the destination includes a network with a smaller MTU. Should the transport protocol that sits on top of IP give IP a packet larger than the local MTU, however, then the source host must fragment it. Fragmentation typically occurs in a router when it receives a datagram that it wants to forward over a network that has an MTU that is smaller than the received datagram. To enable these fragments to be reassembled at the receiving host, they all carry the same identifier in the Ident field. This identifier is chosen by the sending host and is intended to be unique among all the datagrams that might arrive at the destination from this source over some reasonable time period. Since all fragments of the original datagram contain this identifier, the reassembling host will be able to recognize those fragments that go together. Should all the fragments not arrive at the receiving host, the host gives up on the reassembly process and discards the fragments that did arrive. IP does not attempt to recover from missing fragments. To see what this all means, consider what happens when host H1 sends a data- gram to host H8 in the example internet shown in Figure 4.1. Assuming that the MTU 4.1 Simple Internetworking (IP) 243 H1 R1 R2 R3 H8 ETH FDDI PPP IP (376) PPP IP (512) PPP IP (512) (512) ETH IP ETH IP (512)ETH IP (376) IP (1400) IP (1400) R1 R2 R3 Figure 4.4 IP datagrams traversing the sequence of physical networks graphed in Figure 4.1. is 1500 bytes for the two Ethernets, 4500 bytes for the FDDI network, and 532 bytes for the point-to-point network, then a 1420-byte datagram (20-byte IP header plus 1400 bytes of data) sent from H1 makes it across the first Ethernet and the FDDI net- work without fragmentation but must be fragmented into three datagrams at router R2. These three fragments are then forwarded by router R3 across the second Ethernet to the destination host. This situation is illustrated in Figure 4.4. This figure also serves to reinforce two important points: 1 Each fragment is itself a self-contained IP datagram that is transmitted over a sequence of physical networks, independent of the other fragments. 2 Each IP datagram is reencapsulated for each physical network over which it travels. The fragmentation process can be understood in detail by looking at the header fields of each datagram, as is done in Figure 4.5. The unfragmented packet, shown at the top, has 1400 bytes of data and a 20-byte IP header. When the packet arrives at router R2, which has an MTU of 532 bytes, it has to be fragmented. A 532-byte MTU leaves 512 bytes for data after the 20-byte IP header, so the first fragment contains 512 bytes of data. The router sets the M bit in the Flags field (see Figure 4.3), meaning that there are more fragments to follow, and it sets the Offset to 0, since this fragment contains the first part of the original datagram. The data carried in the second fragment starts with the 513th byte of the original data, so the Offset field in this header is set to 64, which is 512 ÷ 8. Why the division by 8? Because the designers of IP decided that fragmentation should always happen on 8-byte boundaries, which means that the Offset field counts 8-byte chunks, not bytes. (We leave it as an exercise for you to figure out why this design decision was made.) The third fragment contains the last 244 4 Internetworking (a) Ident = x Start of header Rest of header 1400 data bytes Offset = 00 (b) Ident = x Start of header Rest of header 512 data bytes Offset = 01 Ident = x Rest of header 512 data bytes Offset = 641 Start of header Ident = x Start of header Rest of header 376 data bytes Offset = 1280 Figure 4.5 Header fields used in IP fragmentation. (a) Unfragmented packet; (b) frag- mented packets. 376 bytes of data, and the offset is now 2 × 512 ÷ 8 = 128. Since this is the last fragment, the M bit is not set. Observe that the fragmentation process is done in such a way that it could be repeated if a fragment arrived at another network with an even smaller MTU. Fragmentation produces smaller, valid IP datagrams that can be readily reassembled 4.1 Simple Internetworking (IP) 245 into the original datagram upon receipt, independent of the order of their arrival. Reassembly is done at the receiving host and not at each router. Implementation We conclude this discussion of IP fragmentation and reassembly by giving a fragment of code that performs reassembly. One reason we give this particular piece of code is that it is representative of a large proportion of networking software—it does little more than tedious and unglamorous bookkeeping. First, we define the key data structure (FragList) that is used to hold the individual fragments that arrive at the destination. Incoming fragments are saved in this data structure until all the fragments in the original datagram have arrived, at which time they are reassembled into a complete datagram and passed up to some higher-level protocol. Note that each element in FragList contains either a fragment or a hole. #define FRAGOFFMASK 0x1fff #define FRAGOFFSET(fragflag) ((fragflag) & FRAGOFFMASK) #define INFINITE_OFFSET 0xffff /* structure to hold the fields that uniquely identify fragments of the same IP datagram */ typedef struct fid { IpHost source; IpHost dest; u_char prot; u_char pad; u_short ident; } FragId; typedef struct hole { u_int first; u_int last; } Hole; #define HOLE 1 #define FRAG 2 /* structure to hold a fragment or a hole */ typedef struct fragif { u_char type; union { Hole hole; Msg frag; }u; struct fragif *next, *prev; } FragInfo; 246 4 Internetworking /* structure to hold all the fragments and holes for a single IP datagram being reassembled */ typedef struct FragList { u_short nholes; FragInfo head; /* dummy header node */ Binding binding; bool gcMark; /* garbage collection flag */ } FragList; The reassembly routine, ipReassemble, takes an incoming datagram (dg) and the IP header for that datagram (hdr) as arguments. The third argument, fragMap, is a Map structure (which supports mapBind, mapRemove, and MapResolve opera- tions) used to efficiently map the incoming datagram into the appropriate FragList. (Recall that the group of fragments that are being reassembled together are uniquely identified by several fields in the IP header, as defined by structure FragId given above.) The actual work done in ipReassemble is straightforward; as stated above, it is mostly bookkeeping. First, the routine extracts the fields from the IP header that uniquely identify the datagram to be reassembled, constructs a key from these fields, and looks this key up in fragMap to find the appropriate FragList. If this is the first fragment for the datagram, a new FragList must be created and initialized. Next, the routine inserts the new fragment into this FragList. This involves comparing the sum of the offset and length of this fragment with the offset of the next fragment in the list. Some of this work is done in subroutine hole create, which is given below. Finally, ipReassemble checks to see if all the holes are filled. If all the fragments are present, it calls the routine msgReassemble to actually reassemble the fragments into a whole datagram and then calls deliver to pass this datagram up the protocol graph to some high-level protocol identified as HLP. ipReassemble(Msg *dg, IpHdr *hdr, Map fragMap) { FragId fragid; FragList *list; FragInfo *fi, *prev; Hole *hole; u_short offset, len; /* extract fragmentation info from header (offset and fragment length) */ offset = FRAGOFFSET(hdr->frag)*8; 4.1 Simple Internetworking (IP) 247 len = hdr->dlen - GET_HLEN(hdr) * 4; /* Create the unique id for this fragment */ bzero((char *)&fragid, sizeof(FragId)); fragid.source = hdr->source; fragid.dest = hdr->dest; fragid.prot = hdr->prot; fragid.ident = hdr->ident; /* find reassembly list for this frag; create one if none exists */ if (mapResolve( fragMap, &fragid, (void **)&list) == FALSE) { /* first fragment of datagram - need new FragList */ list = NEW(FragList); /* insert it into the Map structure */ list->binding = mapBind( fragMap, &fragid, list ); /* initialize list with a single hole spanning the whole datagram */ list->nholes = 1; list->head.next = fi = NEW(FragInfo); fi->next = 0; fi->type = HOLE; fi->u.hole.first = 0; fi->u.hole.last = INFINITE_OFFSET; } /* mark the current FragList as ineligible for garbage collection */ list->gcMark = FALSE; /* walk through the FragList to find the right hole for this frag */ prev = &list->head; for ( fi = prev->next; fi != 0; prev = fi, fi = fi->next ) { if ( fi->type == FRAG ) { continue; } hole = &fi->u.hole; if ( (offset < hole->last) && ((offset + len) > hole->first) ) { /* check to see if frag overlaps previously received frags */ if ( offset < hole->first ) 248 4 Internetworking { /* truncate message from left */ msgStripHdr(dg, hole->first - offset); offset = hole->first; } if ( (offset + len) > hole->last ) { /* truncate message from right */ msgTruncate(dg, hole->last - offset); len = hole->last - offset; } /* now check to see if new hole(s) need to be made */ if (((offset + len) < hole->last) && (hdr->frag & MOREFRAGMENTS)) { /* creating new hole above */ hole_create(prev, fi, (offset+len), hole->last); list->nholes++; } if ( offset > hole->first ) { /* creating new hole below */ hole_create(fi, fi->next, hole->first, (offset)); list->nholes++; } /* change this FragInfo structure to be FRAG */ list->nholes--; fi->type = FRAG; msgSaveCopy(&fi->u.frag, dg); break; } /* if found a hole */ } /* for loop */ /* check to see if we're done, and if so, pass datagram up */ if ( list->nholes == 0 ) { Msg fullMsg; /* now have a full datagram */ for( fi = list->head.next; fi != 0; fi = fi->next ) { msgReassemble(&fullMsg, &fi->u.frag, &fullMsg); } /* get rid of FragList and its Map entry */ 4.1 Simple Internetworking (IP) 249 mapRemove(fragMap, list->binding); ipFreeFragList(list); deliver(HLP, &fullMsg); msgDestroy(&fullMsg); } return SUCCESS; } Subroutine hole create creates a new hole in the fragment list that begins at offset first and continues to offset last. It makes use of the utility NEW, which creates an instance of the given structure. static int hole_create(FragInfo *prev, FragInfo *next, u_int first, u_int last) { FragInfo *fi; /* creating new hole from first to last */ fi = NEW(FragInfo); fi->type = HOLE; fi->u.hole.first = first; fi->u.hole.last = last; fi->next = next; prev->next = fi; } Finally, note that these routines do not capture the entire picture of reassembly. What is not shown is a background process that periodically checks to see if there has been any recent activity on this datagram (it looks at field gcMark), and if not, it deletes the corresponding FragList. IP does not attempt to recover from the situation in which one or more of the fragments does not arrive; it simply gives up and reclaims the memory that was being used for reassembly. One thing to notice from this code is that IP reassembly is far from a simple process. Note, for example, that if a single fragment is lost, the receiver will still attempt to reassemble the datagram, and it will eventually give up and have to garbage- collect the resources that were used to perform the failed reassembly. For this reason, among others, IP fragmentation is generally considered a good thing to avoid. Hosts are now strongly encouraged to perform “path MTU discovery,” a process by which fragmentation is avoided by sending packets that are small enough to traverse the link with the smallest MTU in the path from sender to receiver. 250 4 Internetworking 4.1.3 Global Addresses In the above discussion of the IP service model, we mentioned that one of the things that it provides is an addressing scheme. After all, if you want to be able to send data to any host on any network, there needs to be a way of identifying all the hosts. Thus, we need a global addressing scheme—one in which no two hosts have the same address. Global uniqueness is the first property that should be provided in an addressing scheme. Ethernet addresses are globally unique, but that alone does not suffice for an addressing scheme in a large internetwork. Ethernet addresses are also flat, which means that they have no structure and provide very few clues to routing protocols.2 In contrast, IP addresses are hierarchical, by which we mean that they are made up of several parts that correspond to some sort of hierarchy in the internetwork. Specifically, IP addresses consist of two parts, a network part and a host part. This is a fairly logical structure for an internetwork, which is made up of many interconnected networks. The network part of an IP address identifies the network to which the host is attached; all hosts attached to the same network have the same network part in their IP address. The host part then identifies each host uniquely on that particular network. Thus, in the simple internetwork of Figure 4.1, the addresses of the hosts on network 1, for example, would all have the same network part and different host parts. Note that the routers in Figure 4.1 are attached to two networks. They need to have an address on each network, one for each interface. For example, router R1, which sits between network 2 and network 3, has an IP address on the interface to network 2 that has the same network part as the hosts on network 2, and it has an IP address on the interface to network 3 that has the same network part as the hosts on network 3. Thus, bearing in mind that a router might be implemented as a host with two network interfaces, it is more precise to think of IP addresses as belonging to interfaces than to hosts. Now, what do these hierarchical addresses look like? Unlike some other forms of hierarchical address, the sizes of the two parts are not the same for all addresses. Instead, IP addresses are divided into three different classes, as shown in Figure 4.6, each of which defines different-sized network and host parts. (There are also class D addresses that specify a multicast group, discussed in Section 4.4, and class E addresses that are currently unused.) In all cases, the address is 32 bits long. The class of an IP address is identified in the most significant few bits. If the first bit is 0, it is a class A address. If the first bit is 1 and the second is 0, it is a 2In fact, as we noted, Ethernet addresses do have a structure for the purposes of assignment—the first 24 bits identify the manufacturer—but this provides no useful information to routing protocols since this structure has nothing to do with network topology. 4.1 Simple Internetworking (IP) 251 Network Host 724 0(a) Network Host 14 16 10(b) Network Host 21 8 110(c) Figure 4.6 IP addresses: (a) class A; (b) class B; (c) class C. class B address. If the first two bits are 1 and the third is 0, it is a class C address. Thus, of the approximately 4 billion possible IP addresses, half are class A, one- quarter are class B, and one-eighth are class C. Each class allocates a certain number of bits for the network part of the address and the rest for the host part. Class A networks have 7 bits for the network part and 24 bits for the host part, meaning that there can be only 126 class A networks (the values 0 and 127 are reserved), but each of them can accommodate up to 224 − 2 (about 16 million) hosts (again, there are two reserved values). Class B addresses allocate 14 bits for the network and 16 bits for the host, meaning that each class B network has room for 65,534 hosts. Finally, class C addresses have only 8 bits for the host and 21 for the network part. Therefore, a class C network can have only 256 unique host identifiers, which means only 254 attached hosts (one host identifier, 255, is reserved for broadcast, and 0 is not a valid host number). However, the addressing scheme supports 221 class C networks. On the face of it, this addressing scheme has a lot of flexibility, allowing networks of vastly different sizes to be accommodated fairly efficiently. The original idea was that the Internet would consist of a small number of wide area networks (these would be class A networks), a modest number of site- (campus-) sized networks (these would be class B networks), and a large number of LANs (these would be class C networks). However, as we shall see in Section 4.3, additional flexibility has been needed, and some innovative ways to provide it are now in use. Because one of these techniques actually removes the distinction between address classes, the addressing scheme just described is now known as “classful” addressing to distinguish it from the newer “classless” approach. Before we look at how IP addresses get used, it is helpful to look at some practical matters, such as how you write them down. By convention, IP addresses are written 252 4 Internetworking as four decimal integers separated by dots. Each integer represents the decimal value contained in 1 byte of the address, starting at the most significant. For example, the address of the computer on which this sentence was typed is 171.69.210.245. It is important not to confuse IP addresses with Internet domain names, which are also hierarchical. Domain names tend to be ASCII strings separated by dots, such as cs.princeton.edu. We will be talking about those in Section 9.1. The important thing about IP addresses is that they are what is carried in the headers of IP packets, and it is those addresses that are used in IP routers to make forwarding decisions. 4.1.4 Datagram Forwarding in IP We are now ready to look at the basic mechanism by which IP routers forward data- grams in an internetwork. Recall from Chapter 3 that forwarding is the process of taking a packet from an input and sending it out on the appropriate output, while routing is the process of building up the tables that allow the correct output for a packet to be determined. The discussion here focuses on forwarding; we take up rout- ing in Section 4.2. The main points to bear in mind as we discuss the forwarding of IP datagrams are the following: ■ Every IP datagram contains the IP address of the destination host. ■ The “network part” of an IP address uniquely identifies a single physical network that is part of the larger Internet. ■ All hosts and routers that share the same network part of their address are connected to the same physical network and can thus communicate with each other by sending frames over that network. ■ Every physical network that is part of the Internet has at least one router that, by definition, is also connected to at least one other physical network; this router can exchange packets with hosts or routers on either network. Forwarding IP datagrams can therefore be handled in the following way. A data- gram is sent from a source host to a destination host, possibly passing through sev- eral routers along the way. Any node, whether it is a host or a router, first tries to establish whether it is connected to the same physical network as the destination. To do this, it compares the network part of the destination address with the net- work part of the address of each of its network interfaces. (Hosts normally have only one interface, while routers normally have two or more, since they are typically connected to two or more networks.) If a match occurs, then that means that the destination lies on the same physical network as the interface, and the packet can be 4.1 Simple Internetworking (IP) 253 directly delivered over that network. Section 4.1.5 explains some of the details of this process. If the node is not connected to the same physical network as the destination node, then it needs to send the datagram to a router.In general, each node will have a choice of several routers, and so it needs to pick the best one, or at least one that has a reasonable chance of getting the datagram closer to its destination. The router that it chooses is known as the next hop router. The router finds the correct next hop by consulting its forwarding table. The forwarding table is conceptually just a list of NetworkNum, NextHop pairs. (As we will see below, forwarding tables in practice often contain some additional information related to the next hop.) Normally, there is also a default router that is used if none of the entries in the table match the destination’s network number. For a host, it may be quite acceptable to have a default router and nothing else—this means that all datagrams destined for hosts not on the physical network to which the sending host is attached will be sent out through the default router. We can describe the datagram forwarding algorithm in the following way: if (NetworkNum of destination = NetworkNum of one of my interfaces) then deliver packet to destination over that interface else if (NetworkNum of destination is in my forwarding table) then deliver packet to NextHop router else deliver packet to default router For a host with only one interface and only a default router in its forwarding table, this simplifies to if (NetworkNum of destination = my NetworkNum) then deliver packet to destination directly else deliver packet to default router Let’s see how this works in the example internetwork of Figure 4.1. First, sup- pose that H1 wants to send a datagram to H2. Since they are on the same physical network, H1 and H2 have the same network number in their IP address. Thus, H1 deduces that it can deliver the datagram directly to H2 over the Ethernet. The one issue that needs to be resolved is how H1 finds out the correct Ethernet address for H2—this is the address resolution mechanism described in Section 4.1.5. Now suppose H1 wants to send a datagram to H8. Since these hosts are on different physical networks, they have different network numbers, so H1 deduces that 254 4 Internetworking NetworkNum NextHop 1 R3 2 R1 Table 4.1 Example forwarding table for router R2 in Figure 4.1. it needs to send the datagram to a router. R1 is the only choice—the default router— so H1 sends the datagram over the Ethernet to R1. Similarly, R1 knows that it cannot deliver a datagram directly to H8 because neither of R1’s interfaces is on the same net- work as H8. Suppose R1’s default router is R2; R1 then sends the datagram to R2 over the token ring network. Assuming R2 has the forwarding table shown in Table 4.1, it looks up H8’s network number (network 1) and forwards the datagram to R3. Finally, R3, since it is on the same network as H8, forwards the datagram directly to H8. Note that it is possible to include the infomation about directly connected net- works in the forwarding table. For exam- ple, we would label the network interfaces of router R2 as interface 0 for the point-to- point link (network 4) and interface 1 for the token ring (network 3). Then R2 would have the forwarding table shown in Table 4.2. Thus, for any network number that R2 encounters in a packet, it knows what to do. Either that network is directly connected to R2, in which case the packet can be de- livered to its destination over that network, or the network is reachable via some next Bridges, Switches, and Routers It is easy to become confused about the distinction between bridges, switches, and routers. There is good reason for such confusion, since at some level, they all forward messages from one link to another. One distinction people make is based on layering: Bridges are link- level nodes (they forward frames from one link to another to imple- ment an extended LAN), switches are network-level nodes (they for- ward packets from one link to another to implement a packet- switched network), and routers are internet-level nodes (they forward datagrams from one network to another to implement an internet). In some sense, however, this is an artificial distinction. It is certainly the case that networking compa- nies do not ask the layering police for permission to sell new products that do not fit neatly into one layer or another. 4.1 Simple Internetworking (IP) 255 NetworkNum NextHop 1 R3 2 R1 3 Interface 1 4 Interface 0 Table 4.2 Complete forwarding table for router R2 in Figure 4.1. For example, we have already seen that a multiport bridge is usu- ally called an Ethernet switch or LAN switch. Thus the distinction between bridges and switches has now been largely eroded. For this reason, bridges and switches are often grouped together as “layer 2 devices,” where layer 2 in this context means “above the physical layer, below the internet layer.” There is, however, an im- portant distinction between LAN switches (or bridges) and ATM switches (and other switches that are used in WANs, such as Frame Relay and X.25 switches). LAN switches and bridges depend on the spanning tree algorithm, while WAN switches generally run rout- ing protocols that allow each switch to learn the topology of the whole network. This is an impor- tant distinction because knowing the whole network topology allows hop router that R2 can reach over a net- work to which it is connected. In either case, R2 will use ARP, described below, to find the MAC address of the node to which the packet is to be sent next. The forwarding table used by R2 is simple enough that it could be manually configured. Usually, however, these tables are more complex and would be built up by running a routing protocol such as one of those described in Section 4.2. Also note that, in practice, the network numbers are usually longer (e.g., 128.96). We can now see how hierarchi- cal addressing—splitting the address into network and host parts—has improved the scalability of a large network. Routers now contain forwarding tables that list only a set of network numbers, rather than all the nodes in the network. In our sim- ple example, that meant that R2 could store the information needed to reach all the hosts in the network (of which there were eight) in a four-entry table. Even if there were 100 hosts on each physical net- work, R2 would still only need those same four entries. This is a good first step (al- though by no means the last) in achieving scalability. 256 4 Internetworking ◮ This illustrates one of the most im- portant principles of building scalable net- works: To achieve scalability, you need to reduce the amount of information that is stored in each node and that is exchanged between nodes. The most common way to do that is hierarchical aggregation. IP intro- duces a two-level hierarchy, with networks at the top level and nodes at the bottom level. We have aggregated information by letting routers deal only with reaching the right network; the information that a router needs to deliver a datagram to any node on a given network is represented by a single aggregated piece of information. Router Implementation In Section 3.4 we saw a variety of ways to build a switch, ranging from a general- purpose workstation with a suitable num- ber of network interfaces to some sophisti- cated hardware designs. In general, the same range of options is available for building routers, most of which look something like Figure 4.7. The control processor is respon- sible for running the routing protocols (dis- cussed in Section 4.2) and generally acts as the central point of control of the router. The switching fabric transfers packets from one port to another, just as in a switch, and the ports provide a range of functionality to allow the router to interface to links of various types (e.g., Ethernet, SONET, etc.). A few points are worth noting about router design and how it differs from switch design. First, routers must be designed to handle variable-length packets, a constraint that does not apply to ATM switches but the switches to discriminate among different routes, while in contrast, the spanning tree algorithm locks in a single tree over which messages are forwarded. It is also the case that the spanning tree approach does not scale as well. What about switches and routers? Are they fundamentally the same thing, or are they different in some important way? Here, the distinction is much less clear. For starters, since a single point-to- point link is itself a legitimate network, a router can be used to connect a set of such links. In such a situation, a router looks just like a switch. It just happens to be a switch that forwards IP packets us- ing a datagram forwarding model and IP routing protocols. We’ll see more of this similarity when we consider router implementation at the end of this section. One big difference between an ATM network built from switches and the Internet built from routers is that the Internet is able to accommodate heterogene- ity, whereas ATM consists of ho- mogeneous links. This support for heterogeneity is one of the key rea- sons why the Internet is so widely deployed. 4.1 Simple Internetworking (IP) 257 Switching fabric Control processor Output port Input port Figure 4.7 Block diagram of a router. is certainly applicable to Ethernet or Frame Relay switches. It turns out that many high-performance routers are designed using a switching fabric that is cell based. In such cases the ports must be able to convert variable-length packets into cells and back again. This is very much like the standard ATM segmentation and reassembly (SAR) problem described in Section 3.3.2. Another consequence of the variable length of IP datagrams is that it can be harder to characterize the performance of a router than a switch that forwards only cells. Routers can usually forward a certain number of packets per second, and this implies that the total throughput in bits per second depends on packet size. Router designers generally have to make a choice as to what packet length they will support at line rate. That is, if pps (packets per second) is the rate at which packets arriving on a particular port can be forwarded, and linerate is the physical speed of the port in bits per second, then there will be some packetsize in bits such that packetsize × pps = linerate This is the packet size at which the router can forward at line rate; it is likely to be able to sustain line rate for longer packets but not for shorter packets. Sometimes a designer might decide that the right packet size to support is 40 bytes, since that is the minimum size of an IP packet that has a TCP header attached. Another choice might be the expected average packet size, which can be determined by studying traces of network traffic. For example, measurements of the Internet backbone suggest that the average IP packet is around 300 bytes long. However, such a router would fall behind and perhaps start dropping packets when faced with a long sequence of short 258 4 Internetworking packets, which is statistically likely from time to time and also very possible if the router is subject to an active attack (see Chapter 8). Design decisions of this type depend heavily on cost considerations and the intended application of the router. When it comes to the task of forwarding IP packets, routers can be broadly characterized as having either a centralized or distributed forwarding model. In the centralized model, the IP forwarding algorithm, outlined earlier in this section, is done in a single processing engine that handles the traffic from all ports. In the distributed model, there are several processing engines, perhaps one per port, or more often one per line card, where a line card may serve one or more physical ports. Each model has advantages and disadvantages. All things being equal, a distributed forwarding model should be able to forward more packets per second through the router as a whole because there is more processing power in total. But a distributed model also complicates the software architecture because each forwarding engine typically needs its own copy of the forwarding table, and thus it is necessary for the control processor to ensure that the forwarding tables are updated consistently and in a timely manner. In recent years, there has been considerable interest in the possibility of creating network processors that could be used in the design of routers and other network- ing hardware. A network processor is intended to be a device that is just about as programmable as a standard workstation or PC processor, but that is more highly optimized for networking tasks. For example, a network processor might have in- structions that are particularly well suited to performing lookups on IP addresses or calculating checksums on IP datagrams. One of the interesting and ongoing debates about network processors is whether they can do a better job than the alternatives. For example, given the continuous and remarkable improvements in performance of conventional processors, and the huge industry that drives those improvements, can network processors keep up? And can a device that strives for generality do as good a job as a custom-designed chip that does nothing except, say, IP forwarding? Part of the answer to questions like these depends on what you mean by “do a better job.” For example, there will always be trade-offs to be made between cost of hardware, time to market, performance, and flexibility—the ability to change the features supported by a router after it is built. We will see in the rest of this chapter and in later chapters just how diverse the requirements for router functionality can be. It is safe to assume that a wide range of router designs will exist for the foreseeable future and that network processors will have some role to play. 4.1.5 Address Translation (ARP) In the previous section we talked about how to get IP datagrams to the right physical network, but glossed over the issue of how to get a datagram to a particular host or 4.1 Simple Internetworking (IP) 259 router on that network. The main issue is that IP datagrams contain IP addresses, but the physical interface hardware on the host or router to which you want to send the datagram only understands the addressing scheme of that particular network. Thus, we need to translate the IP address to a link-level address that makes sense on this network (e.g., a 48-bit Ethernet address). We can then encapsulate the IP datagram inside a frame that contains that link-level address and send it either to the ultimate destination or to a router that promises to forward the datagram toward the ultimate destination. One simple way to map an IP address into a physical network address is to encode a host’s physical address in the host part of its IP address. For example, a host with physical address 00100001 01001001 (which has the decimal value 33 in the upper byte and 81 in the lower byte) might be given the IP address 128.96.33.81. While this solution has been used on some networks, it is limited in that the network’s physical addresses can be no more than 16 bits long in this example; they can be only 8 bits long on a class C network. This clearly will not work for 48-bit Ethernet addresses. A more general solution would be for each host to maintain a table of address pairs; that is, the table would map IP addresses into physical addresses. While this table could be centrally managed by a system administrator and then copied to each host on the network, a better approach would be for each host to dynamically learn the contents of the table using the network. This can be accomplished using the Address Resolution Protocol (ARP). The goal of ARP is to enable each host on a network to build up a table of mappings between IP addresses and link-level addresses. Since these mappings may change over time (e.g., because an Ethernet card in a host breaks and is replaced by a new one with a new address), the entries are timed out periodically and removed. This happens on the order of every 15 minutes. The set of mappings currently stored in a host is known as the ARP cache or ARP table. ARP takes advantage of the fact that many link-level network technologies, such as Ethernet and token ring, support broadcast. If a host wants to send an IP datagram to a host (or router) that it knows to be on the same network (i.e., the sending and receiving node have the same IP network number), it first checks for a mapping in the cache. If no mapping is found, it needs to invoke the Address Resolution Protocol over the network. It does this by broadcasting an ARP query onto the network. This query contains the IP address in question (the “target IP address”). Each host receives the query and checks to see if it matches its IP address. If it does match, the host sends a response message that contains its link-layer address back to the originator of the query. The originator adds the information contained in this response to its ARP table. The query message also includes the IP address and link-layer address of the sending host. Thus, when a host broadcasts a query message, each host on the network can learn the sender’s link-level and IP addresses and place that information in its 260 4 Internetworking TargetHardwareAddr (bytes 2–5) TargetProtocolAddr (bytes 0–3) SourceProtocolAddr (bytes 2–3) Hardware type = 1 ProtocolType = 0x0800 SourceHardwareAddr (bytes 4–5) TargetHardwareAddr (bytes 0–1) SourceProtocolAddr (bytes 0–1) HLen = 48 PLen = 32 Operation SourceHardwareAddr (bytes 0–3) 0 8 16 31 Figure 4.8 ARP packet format for mapping IP addresses into Ethernet addresses. ARP table. However, not every host adds this information to its ARP table. If the host already has an entry for that host in its table, it “refreshes” this entry; that is, it resets the length of time until it discards the entry. If that host is the target of the query, then it adds the information about the sender to its table, even if it did not already have an entry for that host. This is because there is a good chance that the source host is about to send it an application-level message, and it may eventually have to send a response or ACK back to the source; it will need the source’s physical address to do this. If a host is not the target and does not already have an entry for the source in its ARP table, then it does not add an entry for the source. This is because there is no reason to believe that this host will ever need the source’s link-level address; there is no need to clutter its ARP table with this information. Figure 4.8 shows the ARP packet format for IP-to-Ethernet address mappings. In fact, ARP can be used for lots of other kinds of mappings—the major differences are in the address sizes. In addition to the IP and link-layer addresses of both sender and target, the packet contains ■ a HardwareType field, which specifies the type of physical network (e.g., Ethernet) ■ a ProtocolType field, which specifies the higher-layer protocol (e.g., IP) ■ HLen (“hardware” address length) and PLen (“protocol” address length) fields, which specify the length of the link-layer address and higher-layer protocol address, respectively ■ an Operation field, which specifies whether this is a request or a response ■ the source and target hardware (Ethernet) and protocol (IP) addresses 4.1 Simple Internetworking (IP) 261 Note that the results of the ARP process can be added as an extra column in a forwarding table like the one in Table 4.1. Thus, for example, when R2 needs to forward a packet to network 2, it not only finds that the next hop is R1, but also finds the MAC address to place on the packet to send it to R1. ATMARP It should be clear that if an ATM network is to operate as part of an IP internetwork, then it too must provide a form of ARP. However, the procedure just described will clearly not work on a simple ATM network, because it depends on the fact that ARP packets can be broadcast to all hosts on a single network. One solution to this problem is to use the LAN emulation procedures described in Section 3.3.5. Since the goal of these procedures is to make an ATM network behave just like a shared-media LAN, which includes support for broadcast, the effect is to reduce ARP to a previously solved problem. There are, however, situations where it may not be desirable to treat an ATM network as an emulated LAN. In particular, LAN emulation can be quite inefficient in a large, wide area ATM network. Recall that in an emulated LAN many packets may need to be sent to the broadcast and unknown server, which then floods those packets to all nodes on the emulated LAN. Clearly, there are limits to how far this can scale. The problem here is that adding broadcast capabilities to an intrinsically nonbroadcast network, while useful in some circumstances, is really overkill if the only reason you need broadcast is to enable address resolution. For this reason, there is a different ARP procedure that may be used in an ATM network and that does not depend on broadcast or LAN emulation. This procedure is known as ATMARP and is part of the Classical IP over ATM model. The reason for calling the model “classical” will become apparent shortly. Like LAN emulation, ATMARP relies on the use of a server to resolve addresses—in this case, it is called an ARP server, and its behavior is described below. A key concept in the Classical IP over ATM model is the logical IP subnet (LIS). The LIS abstraction allows us to take one large ATM network and subdivide it into several smaller subnets. (We define “subnet” precisely in Section 4.3.1, but in this case a subnet behaves much like a single network.) All nodes on the same subnet have the same IP network number. And just as in “classical” IP, two nodes (hosts or routers) that are on the same subnet can communicate directly over the ATM network, whereas two nodes that are on different subnets will have to communicate via one or more routers. An example of an ATM network divided into two LISs appears in Figure 4.9. Note that the IP address of host H1 has a network number of 10, as does the router interface that connects to the left-hand LIS, while H2 has a network number of 12, as does the right-hand interface on the router. That is, H1 and the router connect to the 262 4 Internetworking H2 H1 ATM network 10.0.0.2 10.0.0.1 12.0.0.3 12.0.0.5 R LIS 10 LIS 12 Figure 4.9 Logical IP subnets. same LIS (LIS 10) while H2 is on a different subnet (LIS 12) to which the router also connects. An advantage of the LIS model is that we can connect a large number of hosts and routers to a big ATM network without necessarily giving them all addresses from the same IP network. This may make it easier to manage address assignment, for example, in the case where not all nodes connected to the ATM network are under the control of the same administrative entity. The division of the ATM network into a number of LISs also improves scalability by limiting the number of nodes that must be supported by a single ARP server. The basic job of an ARP server is to enable nodes on a LIS to resolve IP addresses to ATM addresses without using broadcast. Each node in the LIS must be configured with the ATM address of the ARP server, so that it can establish a VC to the server when it boots. Once it has a VC to the server, the node sends a registration message to the ARP server that contains both the IP and ATM addresses of the registering node. Thus the ARP server builds up a complete database of all the IP address, ATM address pairs. Once this is in place, any node that wants to send a packet to some IP address can ask the ARP server to provide the corresponding ATM address. Once this is received, the sending node can use ATM signalling to set up a VC to that ATM address, and then send the packet. Just like conventional ARP, a cache of IP-to-ATM address mappings can be maintained. In addition, the node can keep a VC established to that ATM destination as long as there is enough traffic flowing to justify it, thus avoiding the delay of setting up the VC again when the next packet arrives. An interesting consequence of the Classical IP over ATM model is that two nodes on the same ATM network cannot establish a direct VC between themselves if they are on different subnets. This would violate the rule that communication from one subnet to another must pass through a router. For example, host H1 and host H2 in Figure 4.9 cannot establish a direct VC under the classical model. Instead, each needs 4.1 Simple Internetworking (IP) 263 to have a VC to router R. The simple explanation for this rule is that IP routing is known to work well when that rule is obeyed, as it is in non-ATM networks. New techniques to work around that rule have been developed, but they have introduced considerable complexity and problems of robustness.◮ We have now seen the basic mechanisms that IP provides for dealing with both heterogeneity and scale. On the issue of heterogeneity, IP begins by defining a best- effort service model that makes minimal assumptions about the underlying networks; most notably, this service model is based on unreliable datagrams. IP then makes two important additions to this starting point: (1) a common packet format (fragmenta- tion/reassembly is the mechanism that makes this format work over networks with different MTUs) and (2) a global address space for identifying all hosts (ARP is the mechanism that makes this global address space work over networks with different physical addressing schemes). On the issue of scale, IP uses hierarchical aggregation to reduce the amount of information needed to forward packets. Specifically, IP addresses are partitioned into network and host components, with packets first routed toward the destination network and then delivered to the correct host on that network. 4.1.6 Host Configuration (DHCP) In Section 2.6 we observed that Ethernet addresses are configured into the network adaptor by the manufacturer, and this process is managed in such a way to ensure that these addresses are globally unique. This is clearly a sufficient condition to ensure that any collection of hosts connected to a single Ethernet (including an extended LAN) will have unique addresses. Furthermore, uniqueness is all we ask of Ethernet addresses. IP addresses, by contrast, not only must be unique on a given internetwork, but also must reflect the structure of the internetwork. As noted above, they contain a network part and a host part, and the network part must be the same for all hosts on the same network. Thus, it is not possible for the IP address to be configured once into a host when it is manufactured, since that would imply that the manufacturer knew which hosts were going to end up on which networks, and it would mean that a host, once connected to one network, could never move to another. For this reason, IP addresses need to be reconfigurable. In addition to an IP address, there are some other pieces of information a host needs to have before it can start sending packets. The most notable of these is the address of a default router—the place to which it can send packets whose destination address is not on the same network as the sending host. Most host operating systems provide a way for a system administrator, or even a user, to manually configure the IP information needed by a host. However, there are some obvious drawbacks to such manual configuration. One is that it is simply 264 4 Internetworking a lot of work to configure all the hosts in a large network directly, especially when you consider that such hosts are not reachable over a network until they are config- ured. Even more importantly, the configuration process is very error-prone, since it is necessary to ensure that every host gets the correct network number and that no two hosts receive the same IP address. For these reasons, automated configuration meth- ods are required. The primary method uses a protocol known as the Dynamic Host Configuration Protocol (DHCP). DHCP relies on the existence of a DHCP server that is responsible for providing configuration information to hosts. There is at least one DHCP server for an adminis- trative domain. At the simplest level, the DHCP server can function just as a centralized repository for host configuration information. Consider, for example, the problem of administering addresses in the internetwork of a large company. DHCP saves the net- work administrators from having to walk around to every host in the company with a list of addresses and network map in hand and configuring each host manually. Instead, the configuration information for each host could be stored in the DHCP server and automatically retrieved by each host when it is booted or connected to the network. However, the administrator would still pick the address that each host is to receive; he would just store that in the server. In this model, the configuration information for each host is stored in a table that is indexed by some form of unique client identifier, typically the “hardware address” (e.g., the Ethernet address of its network adaptor). A more sophisticated use of DHCP saves the network admininstrator from even having to assign addresses to individual hosts. In this model, the DHCP server main- tains a pool of available addresses that it hands out to hosts on demand. This consid- erably reduces the amount of configuration an administrator must do, since now it is only necessary to allocate a range of IP addresses (all with the same network number) to each network. Since the goal of DHCP is to minimize the amount of manual configuration required for a host to function, it would rather defeat the purpose if each host had to be configured with the address of a DHCP server. Thus, the first problem faced by DHCP is that of server discovery. To contact a DHCP server, a newly booted or attached host sends a DHCPDISCOVER message to a special IP address (255.255.255.255) that is an IP broad- cast address. This means it will be received by all hosts and routers on that network. (Routers do not forward such packets onto other networks, preventing broadcast to the entire Internet.) In the simplest case, one of these nodes is the DHCP server for the network. The server would then reply to the host that generated the discovery message (all the other nodes would ignore it). However, it is not really desirable to require one DHCP server on every network because this still creates a potentially large number of servers that need to be correctly and consistently configured. Thus, DHCP uses the 4.1 Simple Internetworking (IP) 265 DHCP relay DHCP server Unicast to server Broadcast Host Other networks Figure 4.10 A DHCP relay agent receives a broadcast DHCPDISCOVER message from a host and sends a unicast DHCPDISCOVER message to the DHCP server. Operation HType HLen Hops Xid ciaddr yiaddr siaddr giaddr chaddr (16 bytes) sname (64 bytes) file (128 bytes) options Secs Flags Figure 4.11 DHCP packet format. concept of a relay agent. There is at least one relay agent on each network, and it is configured with just one piece of information: the IP address of the DHCP server. When a relay agent receives a DHCPDISCOVER message, it unicasts it to the DHCP server and awaits the response, which it will then send back to the requesting client. The process of relaying a message from a host to a remote DHCP server is shown in Figure 4.10. Figure 4.11 shows the format of a DHCP message. The message is actually sent using a protocol called UDP (the User Datagram Protocol) that runs over IP. UDP is 266 4 Internetworking discussed in detail in the next chapter, but the only interesting thing it does in this context is to provide a demultiplexing key that says, “This is a DHCP packet.” DHCP is derived from an earlier protocol called BOOTP, and some of the packet fields are thus not strictly relevant to host configuration. When trying to obtain config- uration information, the client puts its hardware address (e.g., its Ethernet address) in the chaddr field. The DHCP server replies by filling in the yiaddr (“your” IP address) field and sending it to the client. Other information such as the default router to be used by this client can be included in the options field. In the case where DHCP dynamically assigns IP addresses to hosts, it is clear that hosts cannot keep addresses indefinitely, as this would eventually cause the server to exhaust its address pool. At the same time, a host cannot be depended upon to give back its address, since it might have crashed, been unplugged from the network, or been turned off. Thus, DHCP allows addresses to be “leased” for some period of time. Once the lease expires, the server is free to return that address to its pool. A host with a leased address clearly needs to renew the lease periodically if in fact it is still connected to the network and functioning correctly.◮ DHCP illustrates an important aspect of scaling: the scaling of network man- agement. While discussions of scaling often focus on keeping the state in network devices from growing too rapidly, it is important to pay attention to growth of net- work management complexity. By allowing network managers to configure a range of IP addresses per network rather than one IP address per host, DHCP improves the manageability of a network. Note that DHCP may also introduce some more complexity into network man- agement, since it makes the binding between physical hosts and IP addresses much more dynamic. This may make the network manager’s job more difficult if, for exam- ple, it becomes necessary to locate a malfunctioning host. 4.1.7 Error Reporting (ICMP) The next issue is how the Internet treats errors. While IP is perfectly willing to drop datagrams when the going gets tough—for example, when a router does not know how to forward the datagram or when one fragment of a datagram fails to arrive at the destination—it does not necessarily fail silently. IP is always configured with a companion protocol, known as the Internet Control Message Protocol (ICMP), that defines a collection of error messages that are sent back to the source host whenever a router or host is unable to process an IP datagram successfully. For example, ICMP defines error messages indicating that the destination host is unreachable (perhaps due to a link failure), that the reassembly process failed, that the TTL had reached 0, that the IP header checksum failed, and so on. 4.1 Simple Internetworking (IP) 267 ICMP also defines a handful of control messages that a router can send back to a source host. One of the most useful control messages, called an ICMP-Redirect, tells the source host that there is a better route to the destination. ICMP-Redirects are used in the following situation. Suppose a host is connected to a network that has two routers attached to it, called R1 and R2, where the host uses R1 as its default router. Should R1 ever receive a datagram from the host, where based on its forwarding table it knows that R2 would have been a better choice for a particular destination address, it sends an ICMP-Redirect back to the host, instructing it to use R2 for all future datagrams addressed to that destination. The host then adds this new route to its forwarding table. 4.1.8 Virtual Networks and Tunnels We conclude our introduction to IP by considering an issue you might not have antici- pated, but one that is becoming increasingly important. Our discussion up to this point has focused on making it possible for nodes on different networks to communicate with each other in an unrestricted way. This is usually the goal in the Internet—everybody wants to be able to send email to everybody, and the creator of a new Web site wants to reach the widest possible audience. However, there are many situations where more controlled connectivity is required. An important example of such a situation is the virtual private network (VPN). The term “VPN” is heavily overused and definitions vary, but intuitively we can define a VPN by considering first the idea of a private network. Corporations with many sites often build private networks by leasing transmission lines from the phone companies and using those lines to interconnect sites. In such a network, communica- tion is restricted to take place only among the sites of that corporation, which is often desirable for security reasons. To make a private network virtual, the leased transmis- sion lines—which are not shared with any other corporations—would be replaced by some sort of shared network. A virtual circuit is a very reasonable replacement for a leased line because it still provides a logical point-to-point connection between the corporation’s sites. For example, if corporation X has a VC from site A to site B, then clearly it can send packets between sites A and B. But there is no way that corpora- tion Y can get its packets delivered to site B without first establishing its own virtual circuit to site B, and the establishment of such a VC can be administratively prevented, thus preventing unwanted connectivity between corporation X and corporation Y. Figure 4.12(a) shows two private networks for two separate corporations. In Figure 4.12(b) they are both migrated to a virtual circuit network. The limited con- nectivity of a real private network is maintained, but since the private networks now share the same transmission facilities and switches we say that two virtual private networks have been created. 268 4 Internetworking Corporation X private network Corporation Y private network Physical links Physical links Virtual circuits (a) (b) C AB KL M KL M C AB Figure 4.12 An example of virtual private networks: (a) two separate private networks; (b) two virtual private networks sharing common switches. In Figure 4.12, a Frame Relay or ATM network is used to provide the con- trolled connectivity among sites. It is also possible to provide a similar function using an IP network—an internetwork—to provide the connectivity. However, we cannot just connect the various corporations’ sites to a single internetwork because that would provide connectivity between corporation X and corporation Y, which we wish to avoid. To solve this problem, we need to introduce a new concept, the IP tunnel. We can think of an IP tunnel as a virtual point-to-point link between a pair of nodes that are actually separated by an arbitrary number of networks. The virtual link is created within the router at the entrance to the tunnel by providing it with the IP address of the router at the far end of the tunnel. Whenever the router at the entrance of the tunnel wants to send a packet over this virtual link, it encapsulates the packet inside an IP datagram. The destination address in the IP header is the address of the 4.1 Simple Internetworking (IP) 269 IP header, Destination = 2.x IP payload IP header, Destination = 10.0.0.1 IP header, Destination = 2.x IP payload IP header, Destination = 2.x IP payload 10.0.0.1 R1 R2InternetworkNetwork 1 Network 2 Figure 4.13 A tunnel through an internetwork. NetworkNum NextHop 1 Interface 0 2 Virtual interface 0 Default Interface 1 Table 4.3 Forwarding table for router R1 in Figure 4.13. router at the far end of the tunnel, while the source address is that of the encapsulating router. In the forwarding table of the router at the entrance to the tunnel, this virtual link looks much like a normal link. Consider, for example, the network in Figure 4.13. A tunnel has been configured from R1 to R2 and assigned a virtual interface number of 0. The forwarding table in R1 might therefore look like Table 4.3. R1 has two physical interfaces. Interface 0 connects to network 1; interface 1 connects to a large internetwork and is thus the default for all traffic that does not match something more specific in the forwarding table. In addition, R1 has a virtual interface, which is the interface to the tunnel. Suppose R1 receives a packet from network 1 that contains an address in network 2. The forwarding table says this packet should be sent out virtual interface 0. In order to send a packet out this interface, the router takes the packet, adds an IP header addressed to R2, and then proceeds to forward the packet as if it had just been received. R2’s address is 10.0.0.1; since the network number of this address is 10, not 1 or 2, a packet destined for R2 will be forwarded out the default interface into the internetwork. 270 4 Internetworking Once the packet leaves R1, it looks to the rest of the world like a normal IP packet destined to R2, and it is forwarded accordingly. All the routers in the internetwork forward it using normal means, until it arrives at R2. When R2 receives the packet, it finds that it carries its own address, so it removes the IP header and looks at the payload of the packet. What it finds is an inner IP packet whose destination address is in network 2. R2 now processes this packet like any other IP packet it receives. Since R2 is directly connected to network 2, it forwards the packet on to that network. Figure 4.13 shows the change in encapsulation of the packet as it moves across the network. While R2 is acting as the endpoint of the tunnel, there is nothing to prevent it from performing the normal functions of a router. For example, it might receive some packets that are not tunneled, but that are addressed to networks it knows how to reach, and it would forward them in the normal way. You might wonder why anyone would want to go to all the trouble of creating a tunnel and changing the encapsulation of a packet as it goes across an internetwork. One reason is security, which we will discuss in more detail in Chapter 8. Supplemented with encryption, a tunnel can become a very private sort of link across a public net- work. Another reason may be that R1 and R2 have some capabilities that are not widely available in the intervening networks, such as multicast routing. By connecting these routers with a tunnel, we can build a virtual network in which all the routers with this capability appear to be directly connected. This in fact is how the MBone (multicast backbone) is built, as we will see in Section 4.4. A third reason to build tunnels is to carry packets from protocols other than IP across an IP network. As long as the routers at either end of the tunnel know how to handle these other protocols, the IP tunnel looks to them like a point-to-point link over which they can send non-IP packets. Tunnels also provide a mechanism by which we can force a packet to be de- livered to a particular place even if its original header—the one that gets encapsulated inside the tunnel header—might suggest that it should go somewhere else. We will see an application of this when we consider mobile hosts in Section 4.2.5. Thus, we see that tunneling is a powerful and quite general technique for building virtual links across internetworks. Tunneling does have its downsides. One is that it increases the length of packets; this might represent a significant waste of bandwidth for short packets. There may also be performance implications for the routers at either end of the tunnel, since they need to do more work than normal forwarding as they add and remove the tunnel header. Finally, there is a management cost for the administrative entity that is responsible for setting up the tunnels and making sure they are correctly handled by the routing protocols. 4.2 Routing 271 4.2 Routing In both this and the previous chapter we have assumed that the switches and routers have enough knowledge of the network topology so they can choose the right port onto which each packet should be output. In the case of virtual circuits, routing is an issue only for the connection request packet; all subsequent packets follow the same path as the request. In datagram networks, including IP networks, routing is an issue for every packet. In either case, a switch or router needs to be able to look at the packet’s destination address and then to determine which of the output ports is the best choice to get the packet to that address. As we saw in Section 3.1.1, the switch makes this decision by consulting a forwarding table. The fundamental problem of routing is, How do switches and routers acquire the information in their forwarding tables?◮ We restate an important distinction, which is often neglected, between forward- ing and routing. Forwarding consists of taking a packet, looking at its destination address, consulting a table, and sending the packet in a direction determined by that table. We saw several examples of forwarding in the preceding section. Routing is the process by which forwarding tables are built. We also note that forwarding is a rel- atively simple and well-defined process performed locally at a node, whereas routing depends on complex distributed algorithms that have continued to evolve throughout the history of networking. While the terms forwarding table and routing table are sometimes used inter- changeably, we will make a distinction between them here. The forwarding table is used when a packet is being forwarded and so must contain enough information to accomplish the forwarding function. This means that a row in the forwarding table contains the mapping from a network number to an outgoing interface and some MAC information, such as the Ethernet address of the next hop. The routing table, on the other hand, is the table that is built up by the routing algorithms as a precursor to building the forwarding table. It generally contains mappings from network numbers to next hops. It may also contain information about how this information was learned, so that the router will be able to decide when it should discard some information. Whether the routing table and forwarding table are actually separate data struc- tures is something of an implementation choice, but there are numerous reasons to keep them separate. For example, the forwarding table needs to be structured to opti- mize the process of looking up a network number when forwarding a packet, while the routing table needs to be optimized for the purpose of calculating changes in topology. In some cases, the forwarding table may even be implemented in specialized hardware, whereas this is rarely if ever done for the routing table. Table 4.4 provides an example of a row from each sort of table. In this case, the routing table tells us that network number 10 is to be reached by a next hop router with the IP address 171.69.245.10, 272 4 Internetworking Network Number NextHop 10 171.69.245.10 (a) Network Number Interface MAC Address 10 if0 8:0:2b:e4:b:1:2 (b) Table 4.4 Example rows from (a) routing and (b) forwarding tables. while the forwarding table contains the information about exactly how to forward a packet to that next hop: Send it out interface number 0 with a MAC address of 8:0:2b:e4:b:1:2. Note that the last piece of information is provided by the Address Resolution Protocol. Before getting into the details of routing, we need to remind ourselves of the key question we should be asking anytime we try to build a mechanism for the Internet: “Does this solution scale?” The answer for the algorithms and protocols described in this section is no. They are designed for networks of fairly modest size—fewer than a hundred nodes, in practice. However, the solutions we describe do serve as a building block for a hierarchical routing infrastructure that is used in the Internet today. Specifically, the protocols described in this section are collectively known as intradomain routing protocols, or interior gateway protocols (IGPs). To understand these terms, we need to define a routing domain: A good working definition is an internetwork in which all the routers are under the same adminstrative control (e.g., a single university campus or the network of a single Internet service provider). The relevance of this definition will become apparent in the next section when we look at interdomain routing protocols. For now, the important thing to keep in mind is that we are considering the problem of routing in the context of small to midsized networks, not for a network the size of the Internet. 4.2.1 Network as a Graph Routing is, in essence, a problem of graph theory. Figure 4.14 shows a graph repre- senting a network. The nodes of the graph, labeled A through F, may be either hosts, switches, routers, or networks. For our initial discussion, we will focus on the case 4.2 Routing 273 4 3 6 21 9 1 1 D A FEB C Figure 4.14 Network represented as a graph. where the nodes are routers. The edges of the graph correspond to the network links. Each edge has an associated cost, which gives some indication of the desirability of sending traffic over that link. A discussion of how edge costs are assigned is given in Section 4.2.4.3 The basic problem of routing is to find the lowest-cost path between any two nodes, where the cost of a path equals the sum of the costs of all the edges that make up the path. For a simple network like the one in Figure 4.14, you could imagine just calculating all the shortest paths and loading them into some nonvolatile storage on each node. Such a static approach has several shortcomings: ■ It does not deal with node or link failures. ■ It does not consider the addition of new nodes or links. ■ It implies that edge costs cannot change, even though we might reasonably wish to temporarily assign a high cost to a link that is heavily loaded. For these reasons, routing is achieved in most practical networks by running routing protocols among the nodes. These protocols provide a distributed, dynamic way to solve the problem of finding the lowest-cost path in the presence of link and node failures and changing edge costs. Note the word “distributed” in the last sentence: It is difficult to make centralized solutions scalable, so all the widely used routing protocols use distributed algorithms. The distributed nature of routing algorithms is one of the main reasons why this has been such a rich field of research and development—there are a lot of challenges in making distributed algorithms work well. For example, distributed algorithms raise the possibility that two routers will at one instant have different ideas about the shortest 3In the example networks (graphs) used throughout this chapter, we use undirected edges and assign each edge a single cost. This is actually a slight simplification. It is more accurate to make the edges directed, which typically means that there would be a pair of edges between each node—one flowing in each direction, and each with its own edge cost. 274 4 Internetworking D G A F E B C Figure 4.15 Distance-vector routing: an example network. path to some destination. In fact, each one may think that the other one is closer to the destination, and decide to send packets to the other one. Clearly, such packets will be stuck in a loop until the discrepancy between the two routers is resolved, and it would be good to resolve it as soon as possible. This is just one example of the type of problem routing protocols must address. To begin our analysis, we assume that the edge costs in the network are known. We will examine the two main classes of routing protocols: distance vector and link state. In Section 4.2.4 we return to the problem of calculating edge costs in a meaningful way. 4.2.2 Distance Vector (RIP) The idea behind the distance-vector algorithm is suggested by its name:4 Each node constructs a one-dimensional array (a vector) containing the “distances” (costs) to all other nodes and distributes that vector to its immediate neighbors. The starting assumption for distance-vector routing is that each node knows the cost of the link to each of its directly connected neighbors. A link that is down is assigned an infinite cost. To see how a distance-vector routing algorithm works, it is easiest to consider an example like the one depicted in Figure 4.15. In this example, the cost of each link is set to 1, so that a least-cost path is simply the one with the fewest hops. (Since all edges have the same cost, we do not show the costs in the graph.) We can represent each node’s knowledge about the distances to all other nodes as a table like the one given in Table 4.5. Note that each node only knows the information in one row of the table (the one that bears its name in the left column). The global view that is presented here is not available at any single point in the network. 4The other common name for this class of algorithm is Bellman-Ford, after its inventors. 4.2 Routing 275 Information Distance to Reach Node Stored at Node A B C D E F G A 0 1 1 ∞ 1 1 ∞ B 1 0 1 ∞ ∞ ∞ ∞ C 1 1 0 1 ∞ ∞ ∞ D ∞ ∞ 1 0 ∞ ∞ 1 E 1 ∞ ∞ ∞ 0 ∞ ∞ F 1 ∞ ∞ ∞ ∞ 0 1 G ∞ ∞ ∞ 1 ∞ 1 0 Table 4.5 Initial distances stored at each node (global view). Destination Cost NextHop B 1 B C 1 C D ∞ — E 1 E F 1 F G ∞ — Table 4.6 Initial routing table at node A. We may consider each row in Table 4.5 as a list of distances from one node to all other nodes, representing the current beliefs of that node. Initially, each node sets a cost of 1 to its directly connected neighbors and ∞ to all other nodes. Thus, A initially believes that it can reach B in one hop and that D is unreachable. The routing table stored at A reflects this set of beliefs and includes the name of the next hop that A would use to reach any reachable node. Initially, then, A’s routing table would look like Table 4.6. The next step in distance-vector routing is that every node sends a message to its directly connected neighbors containing its personal list of distances. For example, 276 4 Internetworking Destination Cost NextHop B 1 B C 1 C D 2 C E 1 E F 1 F G 2 F Table 4.7 Final routing table at node A. node F tells node A that it can reach node G at a cost of 1; A also knows it can reach F at a cost of 1, so it adds these costs to get the cost of reaching G by means of F. This total cost of 2 is less than the current cost of infinity, so A records that it can reach G at a cost of 2 by going through F. Similarly, A learns from C that D can be reached from C at a cost of 1; it adds this to the cost of reaching C (1) and decides that D can be reached via C at a cost of 2, which is better than the old cost of infinity. At the same time, A learns from C that B can be reached from C at a cost of 1, so it concludes that the cost of reaching B via C is 2. Since this is worse than the current cost of reaching B (1), this new information is ignored. At this point, A can update its routing table with costs and next hops for all nodes in the network. The result is shown in Table 4.7. In the absence of any topology changes, it only takes a few exchanges of infor- mation between neighbors before each node has a complete routing table. The pro- cess of getting consistent routing information to all the nodes is called convergence. Table 4.8 shows the final set of costs from each node to all other nodes when routing has converged. We must stress that there is no one node in the network that has all the information in this table—each node only knows about the contents of its own routing table. The beauty of a distributed algorithm like this is that it enables all nodes to achieve a consistent view of the network in the absence of any centralized authority. There are a few details to fill in before our discussion of distance-vector routing is complete. First we note that there are two different circumstances under which a given node decides to send a routing update to its neighbors. One of these circumstances is the periodic update. In this case, each node automatically sends an update message every 4.2 Routing 277 Information Distance to Reach Node Stored at Node A B C D E F G A 0 1 1 2 1 1 2 B 1 0 1 2 2 2 3 C 1 1 0 1 2 2 2 D 2 2 1 0 3 2 1 E 1 2 2 3 0 2 3 F 1 2 2 2 2 0 1 G 2 3 2 1 3 1 0 Table 4.8 Final distances stored at each node (global view). so often, even if nothing has changed. This serves to let the other nodes know that this node is still running. It also makes sure that they keep getting information that they may need if their current routes become unviable. The frequency of these periodic updates varies from protocol to protocol, but it is typically on the order of several seconds to several minutes. The second mechanism, sometimes called a triggered update, happens whenever a node receives an update from one of its neighbors that causes it to change one of the routes in its routing table. That is, whenever a node’s routing table changes, it sends an update to its neighbors, which may lead to a change in their tables, causing them to send an update to their neighbors. Now consider what happens when a link or node fails. The nodes that notice first send new lists of distances to their neighbors, and normally the system settles down fairly quickly to a new state. As to the question of how a node detects a failure, there are a couple of different answers. In one approach, a node continually tests the link to another node by sending a control packet and seeing if it receives an acknowledgment. In another approach, a node determines that the link (or the node at the other end of the link) is down if it does not receive the expected periodic routing update for the last few update cycles. To understand what happens when a node detects a link failure, consider what happens when F detects that its link to G has failed. First, F sets its new distance to G to infinity and passes that information along to A. Since A knows that its 2-hop path to G is through F, A would also set its distance to G to infinity. However, with 278 4 Internetworking the next update from C, A would learn that C has a 2-hop path to G. Thus A would know that it could reach G in 3 hops through C, which is less than infinity, and so A would update its table accordingly. When it advertises this to F, node F would learn that it can reach G at a cost of 4 through A, which is less than infinity, and the system would again become stable. Unfortunately, slightly different circumstances can prevent the network from stabilizing. Suppose, for example, that the link from A to E goes down. In the next round of updates, A advertises a distance of infinity to E, but B and C advertise a distance of 2 to E. Depending on the exact timing of events, the following might happen: Node B, upon hearing that E can be reached in 2 hops from C, concludes that it can reach E in 3 hops and advertises this to A; node A concludes that it can reach E in 4 hops and advertises this to C; node C concludes that it can reach E in 5 hops; and so on. This cycle stops only when the distances reach some number that is large enough to be considered infinite. In the meantime, none of the nodes actually knows that E is unreachable, and the routing tables for the network do not stabilize. This situation is known as the count-to-infinity problem. There are several partial solutions to this problem. The first one is to use some relatively small number as an approximation of infinity. For example, we might decide that the maximum number of hops to get across a certain network is never going to be more than 16, and so we could pick 16 as the value that represents infinity. This at least bounds the amount of time that it takes to count to infinity. Of course, it could also present a problem if our network grew to a point where some nodes were separated by more than 16 hops. One technique to improve the time to stabilize routing is called split horizon. The idea is that when a node sends a routing update to its neighbors, it does not send those routes it learned from each neighbor back to that neighbor. For example, if B has the route (E, 2, A) in its table, then it knows it must have learned this route from A, and so whenever B sends a routing update to A, it does not include the route (E, 2) in that update. In a stronger variation of split horizon, called split horizon with poison reverse, B actually sends that route back to A, but it puts negative information in the route to ensure that A will not eventually use B to get to E. For example, B sends the route (E, ∞) to A. The problem with both of these techniques is that they only work for routing loops that involve two nodes. For larger routing loops, more dras- tic measures are called for. Continuing the above example, if B and C had waited for a while after hearing of the link failure from A before advertising routes to E, they would have found that neither of them really had a route to E. Unfortu- nately, this approach delays the convergence of the protocol; speed of convergence is one of the key advantages of its competitor, link-state routing, the subject of Section 4.2.3. 4.2 Routing 279 Implementation The code that implements this algorithm is very straightforward; we give only some of the basics here. Structure Route defines each entry in the routing table, and constant MAX TTL specifies how long an entry is kept in the table before it is discarded. #define MAX_ROUTES 128 /* maximum size of routing table */ #define MAX_TTL 120 /* time (in seconds) until route expires */ typedef struct { NodeAddr Destination; /* address of destination */ NodeAddr NextHop; /* address of next hop */ int Cost; /* distance metric */ u_short TTL; /* time to live */ } Route; int numRoutes = 0; Route routingTable[MAX_ROUTES]; The routine that updates the local node’s routing table based on a new route is given by mergeRoute. Although not shown, a timer function periodically scans the list of routes in the node’s routing table, decrements the TTL (time to live) field of each route, and discards any routes that have a time to live of 0. Notice, however, that the TTL field is reset to MAX TTL any time the route is reconfirmed by an update message from a neighboring node. void mergeRoute (Route *new) { int i; for (i = 0; i < numRoutes; ++i) { if (new->Destination == routingTable[i].Destination) { if (new->Cost + 1 < routingTable[i].Cost) { /* found a better route: */ break; } else if (new->NextHop == routingTable[i].NextHop) { /* metric for current next hop may have changed: */ break; } else { /* route is uninteresting---just ignore it */ return; } 280 4 Internetworking } } if (i == numRoutes) { /*this is a completely new route; is there room for it?*/ if (numRoutes < MAXROUTES) { ++numRoutes; } else { /* can't fit this route in table so give up */ return; } } routingTable[i] = *new; /* reset TTL */ routingTable[i].TTL = MAX_TTL; /* account for hop to get to next node */ ++routingTable[i].Cost; } Finally, the procedure updateRoutingTable is the main routine that calls merge- Route to incorporate all the routes contained in a routing update that is received from a neighboring node. void updateRoutingTable (Route *newRoute, int numNewRoutes) { int i; for (i=0; i < numNewRoutes; ++i) { mergeRoute(&newRoute[i]); } } Routing Information Protocol (RIP) One of the most widely used routing protocols in IP networks is the Routing Informa- tion Protocol (RIP). Its widespread use is due in no small part to the fact that it was distributed along with the popular Berkeley Software Distribution (BSD) version of Unix, from which many commercial versions of Unix were derived. It is also extremely simple. RIP is the canonical example of a routing protocol built on the distance-vector algorithm just described. Routing protocols in internetworks differ very slightly from the idealized graph model described above. In an internetwork, the goal of the routers is to learn how to 4.2 Routing 281 A C 1 2 3 B D 4 5 6 Figure 4.16 Example network running RIP. Address of net 2 Distance to net 2 Command Must be zero Family of net 2 Address of net 2 Family of net 1 Address of net 1 Address of net 1 Distance to net 1 Version 0 8 16 31 Figure 4.17 RIP packet format. forward packets to various networks. Thus, rather than advertising the cost of reaching other routers, the routers advertise the cost of reaching networks. For example, in Figure 4.16, router C would advertise to router A the fact that it can reach networks 2 and 3 (to which it is directly connected) at a cost of 0; networks 5 and 6 at cost 1; and network 4 at cost 2. We can see evidence of this in the RIP packet format in Figure 4.17. The majority of the packet is taken up with network-address, distance pairs. However, the princi- ples of the routing algorithm are just the same. For example, if router A learns from router B that network X can be reached at a lower cost via B than via the existing next 282 4 Internetworking hop in the routing table, A updates the cost and next hop information for the network number accordingly. RIP is in fact a fairly straightforward implementation of distance-vector routing. Routers running RIP send their advertisements every 30 seconds; a router also sends an update message whenever an update from another router causes it to change its routing table. One point of interest is that it supports multiple address families, not just IP. The network-address part of the advertisements is actually represented as a family, address pair. RIP version 2 (RIPv2) also has some features related to scalability that we will discuss in the next section. As we will see below, it is possible to use a range of different metrics or costs for the links in a routing protocol. RIP takes the simplest approach, with all link costs being equal to 1, just as in our example above. Thus it always tries to find the minimum hop route. Valid distances are 1 through 15, with 16 representing infinity. This also limits RIP to running on fairly small networks—those with no paths longer than 15 hops. 4.2.3 Link State (OSPF) Link-state routing is the second major class of intradomain routing protocol. The starting assumptions for link-state routing are rather similar to those for distance- vector routing. Each node is assumed to be capable of finding out the state of the link to its neighbors (up or down) and the cost of each link. Again, we want to provide each node with enough information to enable it to find the least-cost path to any destination. The basic idea behind link-state protocols is very simple: Every node knows how to reach its directly connected neighbors, and if we make sure that the totality of this knowledge is disseminated to every node, then every node will have enough knowledge of the network to build a complete map of the network. This is clearly a sufficient condition (although not a necessary one) for finding the shortest path to any point in the network. Thus, link-state routing protocols rely on two mechanisms: reliable dissemination of link-state information, and the calculation of routes from the sum of all the accumulated link-state knowledge. Reliable Flooding Reliable flooding is the process of making sure that all the nodes participating in the routing protocol get a copy of the link-state information from all the other nodes. As the term “flooding” suggests, the basic idea is for a node to send its link-state information out on all of its directly connected links, with each node that receives this information forwarding it out on all of its links. This process continues until the information has reached all the nodes in the network. 4.2 Routing 283 More precisely, each node creates an update packet, also called a link-state packet (LSP), that contains the following information: ■ the ID of the node that created the LSP ■ a list of directly connected neighbors of that node, with the cost of the link to each one ■ a sequence number ■ a time to live for this packet The first two items are needed to enable route calculation; the last two are used to make the process of flooding the packet to all nodes reliable. Reliability includes making sure that you have the most recent copy of the information, since there may be multiple, contradictory LSPs from one node traversing the network. Making the flooding reliable has proven to be quite difficult. (For example, an early version of link-state routing used in the ARPANET caused that network to fail in 1981.) Flooding works in the following way. First, the transmission of LSPs between adjacent routers is made reliable using acknowledgments and retransmissions just as in the reliable link-layer protocol described in Section 2.5. However, there are several more steps needed to reliably flood an LSP to all nodes in a network. Consider a node X that receives a copy of an LSP that originated at some other node Y. Note that Y may be any other router in the same routing domain as X. X checks to see if it has already stored a copy of an LSP from Y. If not, it stores the LSP. If it already has a copy, it compares the sequence numbers; if the new LSP has a larger sequence number, it is assumed to be the more recent, and that LSP is stored, replacing the old one. A smaller (or equal) sequence number would imply an LSP older (or not newer) than the one stored, so it would be discarded and no further action would be needed. If the received LSP was the newer one, X then sends a copy of that LSP to all of its neighbors except the neighbor from which the LSP was just received. The fact that the LSP is not sent back to the node from which it was received helps to bring an end to the flooding of an LSP. Since X passes the LSP on to all its neighbors, who then turn around and do the same thing, the most recent copy of the LSP eventually reaches all nodes. Figure 4.18 shows an LSP being flooded in a small network. Each node becomes shaded as it stores the new LSP. In Figure 4.18(a) the LSP arrives at node X, which sends it to neighbors A and C in Figure 4.18(b). A and C do not send it back to X, but send it on to B. Since B receives two identical copies of the LSP, it will accept whichever arrived first and ignore the second as a duplicate. It then passes the LSP on to D, who has no neighbors to flood it to, and the process is complete. 284 4 Internetworking (a) X A CBD (b) X A CBD (c) XA CBD (d) XA CBD Figure 4.18 Flooding of link-state packets. (a) LSP arrives at node X; (b) X floods LSP to A and C; (c) A and C flood LSP to B (but not X); (d) flooding is complete. Just as in RIP, each node generates LSPs under two circumstances. Either the expiry of a periodic timer or a change in topology can cause a node to generate a new LSP. However, the only topology-based reason for a node to generate an LSP is if one of its directly connected links or immediate neighbors has gone down. The failure of a link can be detected in some cases by the link-layer protocol. The demise of a neighbor or loss of connectivity to that neighbor can be detected using periodic “hello” packets. Each node sends these to its immediate neighbors at defined intervals. If a sufficiently long time passes without receipt of a “hello” from a neighbor, the link to that neighbor will be declared down, and a new LSP will be generated to reflect this fact. One of the important design goals of a link-state protocol’s flooding mechanism is that the newest information must be flooded to all nodes as quickly as possible, while old information must be removed from the network and not allowed to circulate. In addition, it is clearly desirable to minimize the total amount of routing traffic that is sent around the network; after all, this is just “overhead” from the perspective of those who actually use the network for their applications. The next few paragraphs describe some of the ways that these goals are accomplished. One easy way to reduce overhead is to avoid generating LSPs unless absolutely necessary. This can be done by using very long timers—often on the order of hours—for the periodic generation of LSPs. Given that the flooding protocol is truly reliable when topology changes, it is safe to assume that messages saying “nothing has changed” do not need to be sent very often. To make sure that old information is replaced by newer information, LSPs carry sequence numbers. Each time a node generates a new LSP, it increments the sequence 4.2 Routing 285 number by 1. Unlike most sequence numbers used in protocols, these sequence numbers are not expected to wrap, so the field needs to be quite large (say, 64 bits). If a node goes down and then comes back up, it starts with a sequence number of 0. If the node was down for a long time, all the old LSPs for that node will have timed out (as described below); otherwise, this node will eventually receive a copy of its own LSP with a higher sequence number, which it can then increment and use as its own sequence number. This will ensure that its new LSP replaces any of its old LSPs left over from before the node went down. LSPs also carry a time to live. This is used to ensure that old link-state information is eventually removed from the network. A node always decrements the TTL of a newly received LSP before flooding it to its neighbors. It also “ages” the LSP while it is stored in the node. When the TTL reaches 0, the node refloods the LSP with a TTL of 0, which is interpreted by all the nodes in the network as a signal to delete that LSP. Route Calculation Once a given node has a copy of the LSP from every other node, it is able to compute a complete map for the topology of the network, and from this map it is able to decide the best route to each destination. The question, then, is exactly how it calculates routes from this information. The solution is based on a well-known algorithm from graph theory—Dijkstra’s shortest-path algorithm. We first define Dijkstra’s algorithm in graph-theoretic terms. Imagine that a node takes all the LSPs it has received and constructs a graphical representation of the net- work, in which N denotes the set of nodes in the graph, l(i, j) denotes the nonnegative cost (weight) associated with the edge between nodes i, j ∈ N, and l(i, j) =∞if no edge connects i and j. In the following description, we let s ∈ N denote this node, that is, the node executing the algorithm to find the shortest path to all the other nodes in N. Also, the algorithm maintains the following two variables: M denotes the set of nodes incorporated so far by the algorithm, and C(n) denotes the cost of the path from s to each node n. Given these definitions, the algorithm is defined as follows: M ={s}for each n in N −{s}C(n) = l(s, n) while (N = M) M = M ∪{w} such that C(w) is the minimum for all w in (N − M) for each n in (N − M) C(n) = MIN(C(n), C(w) + l(w, n)) Basically, the algorithm works as follows. We start with M containing this node s and then initialize the table of costs (the C(n)s) to other nodes using the known costs to 286 4 Internetworking directly connected nodes. We then look for the node that is reachable at the lowest cost (w) and add it to M. Finally, we update the table of costs by considering the cost of reaching nodes through w. In the last line of the algorithm, we choose a new route to node n that goes through node w if the total cost of going from the source to w and then following the link from w to n is less than the old route we had to n. This procedure is repeated until all nodes are incorporated in M. In practice, each switch computes its routing table directly from the LSPs it has collected using a realization of Dijkstra’s algorithm called the forward search algo- rithm. Specifically, each switch maintains two lists, known as Tentative and Confirmed. Each of these lists contains a set of entries of the form (Destination, Cost, NextHop). The algorithm works as follows: 1 Initialize the Confirmed list with an entry for myself; this entry has a cost of 0. 2 For the node just added to the Confirmed list in the previous step, call it node Next, select its LSP. 3 For each neighbor (Neighbor)ofNext, calculate the cost (Cost) to reach this Neighbor as the sum of the cost from myself to Next and from Next to Neighbor. (a) If Neighbor is currently on neither the Confirmed nor the Tentative list, then add (Neighbor, Cost, NextHop) to the Tentative list, where NextHop is the direction I go to reach Next. (b) If Neighbor is currently on the Tentative list, and the Cost is less than the cur- rently listed cost for Neighbor, then replace the current entry with (Neighbor, Cost, NextHop), where NextHop is the direction I go to reach Next. 4 If the Tentative list is empty, stop. Otherwise, pick the entry from the Tentative list with the lowest cost, move it to the Confirmed list, and return to step 2. This will become a lot easier to understand when we look at an example. Consider the network depicted in Figure 4.19. Note that, unlike our previous example, this D A B C 5 3 2 11 10 Figure 4.19 Link-state routing: an example network. 4.2 Routing 287 Step Confirmed Tentative Comments 1 (D,0,-) Since D is the only new member of the confirmed list, look at its LSP. 2 (D,0,-) (B,11,B) (C,2,C) D’s LSP says we can reach B through B at cost 11, which is better than anything else on either list, so put it on Tentative list; same for C. 3 (D,0,-) (C,2,C) (B,11,B) Put lowest-cost member of Tentative (C) onto Confirmed list. Next, examine LSP of newly confirmed member (C). 4 (D,0,-) (C,2,C) (B,5,C) (A,12,C) Cost to reach B through C is 5, so replace (B,11,B). C’s LSP tells us that we can reach A at cost 12. 5 (D,0,-) (C,2,C) (B,5,C) (A,12,C) Move lowest-cost member of Tentative (B) to Confirmed, then look at its LSP. 6 (D,0,-) (C,2,C) (B,5,C) (A,10,C) Since we can reach A at cost 5 through B, replace the Tentative entry. 7 (D,0,-) (C,2,C) (B,5,C) (A,10,C) Move lowest-cost member of Tentative (A) to Confirmed, and we are all done. Table 4.9 Steps for building routing table for node D (Figure 4.19). network has a range of different edge costs. Table 4.9 traces the steps for building the routing table for node D. We denote the two outputs of D by using the names of the nodes to which they connect, B and C. Note the way the algorithm seems to head off on false leads (like the 11-unit cost path to B that was the first addition to the Tentative list) but ends up with the least-cost paths to all nodes. 288 4 Internetworking The link-state routing algorithm has many nice properties: It has been proven to stabilize quickly, it does not generate much traffic, and it responds rapidly to topology changes or node failures. On the downside, the amount of information stored at each node (one LSP for every other node in the network) can be quite large. This is one of the fundamental problems of routing and is an instance of the more general prob- lem of scalability. Some solutions to both the specific problem (the amount of storage potentially required at each node) and the general problem (scalability) will be dis- cussed in the next section.◮ Thus, the difference between the distance-vector and link-state algorithms can be summarized as follows. In distance vector, each node talks only to its directly connected neighbors, but it tells them everything it has learned (i.e., distance to all nodes). In link state, each node talks to all other nodes, but it tells them only what it knows for sure (i.e., only the state of its directly connected links). The Open Shortest Path First Protocol (OSPF) One of the most widely used link-state routing protocols is OSPF. The first word, “Open,” refers to the fact that it is an open, nonproprietary standard, created under the auspices of the IETF. The “SPF” part comes from an alternative name for link- state routing. OSPF adds quite a number of features to the basic link-state algorithm described above, including the following: ■ Authentication of routing messages: This is a nice feature, since it is all too common for some misconfigured host to decide that it can reach every host in the universe at a cost of 0. When the host advertises this fact, every router in the surrounding neighborhood updates its forwarding tables to point to that host, and said host receives a vast amount of data that, in reality, it has no idea what to do with. It typically drops it all, bringing the network to a halt. Such disasters can be averted in many cases by requiring routing updates to be authenticated. Early versions of OSPF used a simple 8-byte password for authentication. This is not a strong enough form of authentication to prevent dedicated malicious users, but it alleviates many problems caused by misconfiguration. (A similar form of authentication was added to RIP in version 2.) Strong cryptographic authentication of the sort discussed in Section 8.2.1 was later added. ■ Additional hierarchy: Hierarchy is one of the fundamental tools used to make systems more scalable. OSPF introduces another layer of hierarchy into rout- ing by allowing a domain to be partitioned into areas. This means that a router within a domain does not necessarily need to know how to reach every network within that domain—it may be sufficient for it to know only how to 4.2 Routing 289 Authentication Version Type Message length Checksum Authentication type SourceAddr AreaId 0 8 16 31 Figure 4.20 OSPF header format. get to the right area. Thus, there is a reduction in the amount of information that must be transmitted to and stored in each node. We examine areas in detail in Section 4.3.4. ■ Load balancing: OSPF allows multiple routes to the same place to be assign- ed the same cost and will cause traffic to be distributed evenly over those routes. There are several different types of OSPF messages, but all begin with the same header, as shown in Figure 4.20. The Version field is currently set to 2, and the Type field may take the values 1 through 5. The SourceAddr identifies the sender of the message, and the AreaId is a 32-bit identifier of the area in which the node is located. The entire packet, except the authentication data, is protected by a 16-bit checksum using the same algorithm as the IP header (see Section 2.4). The Authentication type is 0 if no authentication is used; otherwise it may be 1, implying a simple password is used, or 2, which indicates that a cryptographic authentication checksum, of the sort described in Section 8.2.1, is used. In the latter cases the Authentication field carries the password or cryptographic checksum. Of the five OSPF message types, type 1 is the “hello” message, which a router sends to its peers to notify them that it is still alive and connected as described above. The remaining types are used to request, send, and acknowledge the receipt of link- state messages. The basic building block of link-state messages in OSPF is known as the link-state advertisement (LSA). One message may contain many LSAs. We provide a few details of the LSA here. Like any internetwork routing protocol, OSPF must provide information about how to reach networks. Thus, OSPF must provide a little more information than the simple graph-based protocol described above. Specifically, a router running OSPF may generate link-state packets that advertise one or more of the networks that are directly 290 4 Internetworking LS Age Options Type=1 0 Flags 0 Number of links Link type Num_TOS Metric Link-state ID Advertising router LS sequence number Link ID Link data Optional TOS information More links LS checksum Length Figure 4.21 OSPF link-state advertisement. connected to that router. In addition, a router that is connected to another router by some link must advertise the cost of reaching that router over the link. These two types of advertisements are necessary to enable all the routers in a domain to determine the cost of reaching all networks in that domain and the appropriate next hop for each network. Figure 4.21 shows the packet format for a “type 1” link-state advertisement. Type 1 LSAs advertise the cost of links between routers. Type 2 LSAs are used to advertise networks to which the advertising router is connected, while other types are used to support additional hierarchy as described in the next section. Many fields in the LSA should be familiar from the preceding discussion. The LS Age is the equivalent of a time to live, except that it counts up and the LSA expires when the age reaches a defined maximum value. The Type field tells us that this is a type 1 LSA. In a type 1 LSA, the Link-state ID and the Advertising router field are identical. Each carries a 32-bit identifier for the router that created this LSA. While a number of assignment strategies may be used to assign this ID, it is essential that it be unique in the routing domain and that a given router consistently uses the same router ID. One way to pick a router ID that meets these requirements would be to pick the lowest IP address among all the IP addresses assigned to that router. (Recall that a router may have a different IP address on each of its interfaces.) The LS sequence number is used exactly as described above, to detect old or duplicate LSAs. The LS checksum is similar to others we have seen in Section 2.4 and in other protocols; it is of course used to verify that data has not been corrupted. It covers all fields in the packet except LS Age, so that it is not necessary to recompute a checksum every time LS Age is incremented. Length is the length in bytes of the complete LSA. 4.2 Routing 291 Now we get to the actual link-state information. This is made a little complicated by the presence of TOS (type of service) information. Ignoring that for a moment, each link in the LSA is represented by a Link ID, some Link Data, and a metric. The first two of these fields identify the link; a common way to do this would be to use the router ID of the router at the far end of the link as the Link ID, and then use the Link Data to disam- biguate among multiple parallel links if necessary. The metric is of course the cost of the link. Type tells us something about the link, for example, if it is a point-to-point link. The TOS information is present to allow OSPF to choose different routes for IP packets based on the value in their TOS field. Instead of assigning a single metric to a link, it is possible to assign different metrics depending on the TOS value of the data. For example, if we had a link in our network that was very good for delay-sensitive traffic, we could give it a low metric for the TOS value representing low delay and a high metric for everything else. OSPF would then pick a different shortest path for those packets that had their TOS field set to that value. It is worth noting that, at the time of writing, this capability has not been widely deployed.5 4.2.4 Metrics The preceding discussion assumes that link costs, or metrics, are known when we execute the routing algorithm. In this section, we look at some ways to calculate link costs that have proven effective in practice. One example that we have seen already, which is quite reasonable and very simple, is to assign a cost of 1 to all links—the least-cost route will then be the one with the fewest hops. Such an approach has several drawbacks, however. First, it does not distinguish between links on a latency basis. Thus, a satellite link with 250-ms latency looks just as attractive to the routing protocol as a terrestrial link with 1-ms latency. Second, it does not distinguish between routes on a capacity basis, making a 9.6-Kbps link look just as good as a 45-Mbps link. Finally, it does not distinguish between links based on their current load, making it impossible to route around overloaded links. It turns out that this last problem is the hardest because you are trying to capture the complex and dynamic characteristics of a link in a single scalar cost. The ARPANET was the testing ground for a number of different approaches to link-cost calculation. (It was also the place where the superior stability of link state over distance-vector routing was demonstrated; the original mechanism used distance vector while the later version used link state.) The following discussion traces the evolution of the ARPANET routing metric and, in so doing, explores the subtle aspects of the problem. 5Note also that the meaning of the TOS field has changed since the OSPF specification was written. This topic is discussed in Section 6.5.3. 292 4 Internetworking The original ARPANET routing metric measured the number of packets that were queued waiting to be transmitted on each link, meaning that a link with 10 packets queued waiting to be transmitted was assigned a larger cost weight than a link with 5 packets queued for transmission. Using queue length as a routing metric did not work well, however, since queue length is an artificial measure of load—it moves packets toward the shortest queue rather than toward the destination, a situation all too familiar to those of us who hop from line to line at the grocery store. Stated more precisely, the original ARPANET routing mechanism suffered from the fact that it did not take either the bandwidth or the latency of the link into consideration. A second version of the ARPANET routing algorithm, sometimes called the “new routing mechanism,” took both link bandwidth and latency into consideration and used delay, rather than just queue length, as a measure of load. This was done as follows. First, each incoming packet was timestamped with its time of arrival at the router (ArrivalTime); its departure time from the router (DepartTime) was also recorded. Second, when the link-level ACK was received from the other side, the node computed the delay for that packet as Delay = (DepartTime − ArrivalTime) + TransmissionTime + Latency where TransmissionTime and Latency were statically defined for the link and captured the link’s bandwidth and latency, respectively. Notice that in this case, DepartTime – ArrivalTime represents the amount of time the packet was delayed (queued) in the node due to load. If the ACK did not arrive, but instead the packet timed out, then Depart- Time was reset to the time the packet was retransmitted. In this case, DepartTime – ArrivalTime captures the reliability of the link—the more frequent the retransmission of packets, the less reliable the link, and the more we want to avoid it. Finally, the weight assigned to each link was derived from the average delay experienced by the packets recently sent over that link. Although an improvement over the original mechanism, this approach also had a lot of problems. Under light load, it worked reasonably well, since the two static factors of delay dominated the cost. Under heavy load, however, a congested link would start to advertise a very high cost. This caused all the traffic to move off that link, leaving it idle, so then it would advertise a low cost, thereby attracting back all the traffic, and so on. The effect of this instability was that, under heavy load, many links would in fact spend a great deal of time being idle, which is the last thing you want under heavy load. Another problem was that the range of link values was much too large. For example, a heavily loaded 9.6-Kbps link could look 127 times more costly than a lightly loaded 56-Kbps link. This means that the routing algorithm would choose a path with 126 hops of lightly loaded 56-Kbps links in preference to a 1-hop 9.6-Kbps path. 4.2 Routing 293 While shedding some traffic from an overloaded line is a good idea, making it look so unattractive that it loses all its traffic is excessive. Using 126 hops when 1 hop will do is in general a bad use of network resources. Also, satellite links were unduly penalized, so that an idle 56-Kbps satellite link looked considerably more costly than an idle 9.6-Kbps terrestrial link, even though the former would give better performance for high-bandwidth applications. A third approach, called the “revised ARPANET routing metric,” addressed these problems. The major changes were to compress the dynamic range of the metric considerably, to account for the link type, and to smooth the variation of the metric with time. The smoothing was achieved by several mechanisms. First, the delay measure- ment was transformed to a link utilization, and this number was averaged with the last reported utilization to suppress sudden changes. Second, there was a hard limit on how much the metric could change from one measurement cycle to the next. By smoothing the changes in the cost, the likelihood that all nodes would abandon a route at once is greatly reduced. The compression of the dynamic range was achieved by feeding the measured utilization, the link type, and the link speed into a function that is shown graphically in Figure 4.22. Observe the following: 225 New metric (routing units) 140 90 75 60 30 25% 50% 75% 100% Utilization 9.6-Kbps satellite link 9.6-Kbps terrestrial link 56-Kbps satellite link 56-Kbps terrestrial link Figure 4.22 Revised ARPANET routing metric versus link utilization. 294 4 Internetworking ■ A highly loaded link never shows a cost of more than three times its cost when idle. ■ The most expensive link is only seven times the cost of the least ex- pensive. ■ A high-speed satellite link is more attractive than a low-speed terres- trial link. ■ Cost is a function of link uti- lization only at moderate to high loads. All these factors mean that a link is much less likely to be universally abandoned, since a threefold increase in cost is likely to make the link unattractive for some paths while letting it remain the best choice for others. The slopes, offsets, and breakpoints for the curves in Figure 4.22 were arrived at by a great deal of trial and error, and they were carefully tuned to provide good per- formance. There is one final issue related to cal- culating edge weights—the frequency with which each node calculates the weights on its links. There are two things to keep in mind. First, none of the metrics are instan- taneous. That is, whether a node is measur- ing queue length, delay, or utilization, it is actually computing an average over a pe- riod of time. Second, just because a met- ric changes does not mean that the node sends out an update message. In practice, updates are sent only when the change to an edge weight is larger than some threshold. Monitoring Routing Behavior Given the complexity of routing packets through a network of the scale of the Internet, we might won- der how well the system works. We know it works some of the time because we are able to connect to sites all over the world. We suspect it doesn’t work all the time, though, because sometimes we are unable to connect to certain sites. The real problem is determining what part of the system is at fault when our connections fail: Has some routing machinery failed to work properly, is the remote server too busy, or has some link or machine simply gone down? This is really an issue of net- work management, and while there are tools that system administra- tors use to keep tabs on their own networks—for example, see the Simple Network Management Protocol (SNMP) described in Sec- tion 9.2.3—it is a largely unre- solved problem for the Internet as a whole. In fact, the Internet has grown so large and complex that, even though it is constructed from a collection of man-made, largely de- terministic parts, we have come to view it almost as a living organism 4.2 Routing 295 or natural phenomenon that is to be studied. That is, we try to under- stand the Internet’s dynamic behav- ior by performing experiments on it and proposing models that explain our observations. An excellent example of this kind of study has been conducted by Vern Paxson. Paxson used the Unix traceroute tool to study 40,000 end-to-end routes between 37 Internet sites in 1995. He was attempting to answer questions about how routes fail, how stable routes are over time, and whether or not they are symmetric. Among other things, Paxson found that the likelihood of a user encountering a serious end-to-end routing problem was 1 in 30, and that such problems usually lasted about 30 seconds. He also found that two-thirds of the Internet’s routes persisted for days or weeks, and that about one-third of the time the route used to get from host A to host B included at least one different routing domain than the route used to get from host B to host A. Paxson’s over- all conclusion was that Internet routing was becoming less and less predictable over time. 4.2.5 Routing for Mobile Hosts Looking back over the preceding discussion of how IP addressing and routing works, you might notice that there is an implicit assumption about the mobility of hosts, or rather the lack of it. A host’s address con- sists of a network number and a host part, and the network number tells us which net- work the host is attached to. IP routing al- gorithms tell the routers how to get packets to the correct network, thus enhancing the scalability of the routing system by keeping host-specific information out of the routers. So what would happen if a host were dis- connected from one network and connected to another? If we didn’t change the IP address of the host, then it would become unreachable. Any packet destined for this host would be sent to the network that has the appropriate network number, but when the router(s) on that network tried to deliver the packet to the host, the host would not be there to receive it. The obvious solution to this problem is to provide the host with a new address when it attaches to a new network. Tech- niques such as DHCP (described in Sec- tion 4.1.6) can make this a relatively simple process. In many situations this solution is adequate, but in others it is not. For exam- ple, suppose that a user of a PC equipped with a wireless network interface is running some application while she roams the coun- tryside. The PC might detach itself from one network and attach to another with some frequency, but the user would want to be oblivious to this. In particular, the 296 4 Internetworking applications that were running when the PC was attached to network A should con- tinue to run without interruption when it attaches to network B. If the PC simply changes its IP address in the middle of running the application, the application cannot simply keep working, because the remote end has no way of knowing that it must now send the packets to a new IP address. Ideally, we want the movement of the PC to be transparent to the remote application. The procedures that are designed to address this problem are usually referred to as “Mobile IP” (which is also the name of the IETF working group that defined them). The Mobile IP working group made some important design decisions at the outset. In particular, it was a requirement that the solution would work without any changes to the software of nonmobile hosts or the majority of routers in the Internet. This sort of approach is frequently adopted in the Internet. Any new technology that requires a majority of routers or hosts to be modified before it can work is likely to face an uphill battle for acceptance. While the majority of routers remain unchanged, mobility support does require some new functionality in at least one router, known as the home agent of the mo- bile node. This router is located on the “home” network of the mobile host. The mobile host is assumed to have a permanent IP address, called its home address, which has a network number equal to that of the home network, and thus of the home agent. This is the address that will be used by other hosts when they send packets to the mobile host; since it does not change, it can be used by long-lived applications as the host roams. In many cases, a second router with enhanced functionality, the foreign agent, is also required. This router is located on a network to which the mobile node attaches itself when it is away from its home network. We will consider first the operation of Mobile IP when a foreign agent is used. An example network with both home and foreign agents is shown in Figure 4.23. Both home and foreign agents periodically announce their presence on the net- works to which they are attached using agent advertisement messages. A mobile host Internetwork Foreign agent (12.0.0.6) Mobile host (10.0.0.9) Home agent (10.0.0.3) Home network (network 10) Sending host Figure 4.23 Mobile host and mobility agents. 4.2 Routing 297 may also solicit an advertisement when it attaches to a new network. The advertise- ment by the home agent enables a mobile host to learn the address of its home agent before it leaves its home network. When the mobile host attaches to a foreign network, it hears an advertisement from a foreign agent and registers with the agent, provid- ing the address of its home agent. The foreign agent then contacts the home agent, providing a care-of address. This is usually the IP address of the foreign agent. At this point, we can see that any host that tries to send a packet to the mobile host will send it with a destination address equal to the home address of that node. Normal IP forwarding will cause that packet to arrive on the home network of the mobile node, on which the home agent is sitting. Thus, we can divide the problem of delivering the packet to the mobile node into three parts: 1 How does the home agent intercept a packet that is destined for the mobile node? 2 How does the home agent then deliver the packet to the foreign agent? 3 How does the foreign agent deliver the packet to the mobile node? The first problem might look easy if you just look at Figure 4.23, in which the home agent is clearly the only path between the sending host and the home network, and thus must receive packets that are destined to the mobile node. But what if the sending node were on network 10, or what if there were another router connected to network 10 that tried to deliver the packet without its passing through the home agent? To address this problem, the home agent actually impersonates the mobile node, using a technique called “proxy ARP.” This works just like ARP as described in Section 4.1.5, except that the home agent inserts the IP address of the mobile node, rather than its own, in the ARP messages. It uses its own hardware address, so that all the nodes on the same network learn to associate the hardware address of the home agent with the IP address of the mobile node. One subtle aspect of this process is the fact that ARP information may be cached in other nodes on the network. To make sure that these caches are invalidated in a timely way, the home agent issues an ARP message as soon as the mobile node registers with a foreign agent. Because the ARP message is not a response to a normal ARP request, it is termed a “gratuitous ARP.” The second problem is the delivery of the intercepted packet to the foreign agent. Here we use the tunneling technique described in Section 4.1.8. The home agent simply “wraps” the packet inside an IP header that is destined for the foreign agent and transmits it into the internetwork. All the intervening routers just see an IP packet destined for the IP address of the foreign agent. Another way of looking at this is that an IP tunnel is established between the home agent and the foreign agent, and the home agent just drops packets destined for the mobile node into that tunnel. 298 4 Internetworking When a packet finally arrives at the foreign agent, it strips the extra IP header and finds inside an IP packet destined for the mobile node. Clearly, the foreign agent cannot treat this like any old IP packet because this would cause it to send it back to the home network. Instead, it has to recognize the address as that of a regis- tered mobile node. It then delivers the packet to the hardware address of the mobile node (e.g., its Ethernet address), which was learned as part of the registration process. One observation that can be made about these procedures is that it is possible for the foreign agent and the mobile node to be in the same box; that is, a mobile node can perform the foreign agent function itself. To make this work, however, the mobile node must be able to dynamically acquire an IP address that is located in the address space of the foreign network. This address will then be used as the care-of address. In our example, this would have to be an address with a network number of 12. We have already seen one way in which a host can dynamically acquire a correct IP address, using DHCP (Section 4.1.6). This approach has the desirable feature of allowing mobile nodes to attach to networks that don’t have foreign agents; thus, mobility can be achieved with only the addition of a home agent and some new software on the mobile node (assuming DHCP is used on the foreign network). What about traffic in the other direction (i.e., from mobile node to fixed node)? This turns out to be much easier. The mobile node just puts the IP address of the fixed node in the destination field of its IP packets, while putting its permanent address in the source field, and the packets are forwarded to the fixed node using normal means. Of course, if both nodes in a conversation are mobile, then the procedures described above are used in each direction. Route Optimization in Mobile IP There is one significant drawback to the above approach, which may be familiar to users of cellular telephones. The route from sending node to mobile node can be sig- nificantly suboptimal. One of the most extreme examples is when a mobile node and the sending node are on the same network, but the home network for the mobile node is on the far side of the Internet. The sending node addresses all packets to the home network; they traverse the Internet to reach the home agent, which then tunnels them back across the Internet to reach the foreign agent. Clearly, it would be nice if the sending node could find out that the mobile node is actually on the same network and deliver the packet directly. In the more general case, the goal is to deliver packets as directly as possible from sending node to mobile node without passing through a home agent. This is sometimes referred to as the “triangle routing problem” since the path from sender to mobile node via home agent takes two sides of a triangle, rather than the third side that is the direct path. The basic idea behind the solution to triangle routing is to let the sending node know the care-of address of the mobile node. The sending node can then create its 4.3 Global Internet 299 own tunnel to the foreign agent. This is treated as an optimization of the process just described. If the sender has been equipped with the necessary software to learn the care-of address and create its own tunnel, then the route can be optimized; if not, packets just follow the suboptimal route. When a home agent sees a packet destined for one of the mobile nodes that it supports, it can deduce that the sender is not using the optimal route. Therefore, it sends a “binding update” message back to the source, in addition to forwarding the data packet to the foreign agent. The source, if capable, uses this binding update to create an entry in a “binding cache,” which consists of a list of mappings from mobile node addresses to care-of addresses. The next time this source has a data packet to send to that mobile node, it will find the binding in the cache and can tunnel the packet directly to the foreign agent. There is an obvious problem with this scheme, which is that the binding cache may become out-of-date if the mobile host moves to a new network. If an out-of-date cache entry is used, the foreign agent will receive tunneled packets for a mobile node that is no longer registered on its network. In this case, it sends a “binding warning” message back to the sender to tell it to stop using this cache entry. This scheme works only in the case where the foreign agent is not the mobile node itself, however. For this reason, cache entries need to be deleted after some period of time; the exact amount is specified in the binding update message. Mobile routing provides some interesting security challenges. For example, an attacker wishing to intercept the packets destined to some other node in an inter- network could contact the home agent for that node and announce itself as the new foreign agent for the node. Thus it is clear that some authentication mechanisms are required. We discuss such mechanisms in Chapter 8. Finally, we note that there are many open issues in mobile networking. For example, the security and performance aspects of mobile networks might require rout- ing algorithms to take account of several factors when finding a route to a mobile host; for example, it might be desirable to find a route that doesn’t pass through some un- trusted network. There is also the problem of “ad hoc” mobile networks—enabling a group of mobile nodes to form a network in the absence of any fixed nodes. These continue to be areas of active research. 4.3 Global Internet At this point, we have seen how to connect a heterogeneous collection of networks to create an internetwork and how to use the simple hierarchy of the IP address to make routing in an internet somewhat scalable. We say “somewhat” scalable because even though each router does not need to know about all the hosts connected to the internet, 300 4 Internetworking NSFNET backboneStanford BARRNET regional Berkeley PARC NCAR UA UNM Westnet regional UNL KU ISU MidNet regional… Figure 4.24 The tree structure of the Internet in 1990. it does, in the model described so far, need to know about all the networks connected to the internet. Today’s Internet has tens of thousands of networks connected to it. Routing protocols such as those we have just discussed do not scale to those kinds of numbers. This section looks at a variety of techniques that greatly improve scalability and that have enabled the Internet to grow as far as it has. Before getting to these techniques, we need to have a general picture in our heads of what the global Internet looks like. It is not just a random interconnection of Ethernets, but instead it takes on a shape that reflects the fact that it interconnects many different organizations. Figure 4.24 gives a simple depiction of the state of the Internet in 1990. Since that time, the Internet’s topology has grown much more complex than this figure suggests—we present a more accurate picture of the current Internet in Section 4.3.3 and Figure 4.29—but this picture will do for now. One of the salient features of this topology is that it consists of “end user” sites (e.g., Stanford University) that connect to “service provider” networks (e.g., BARR- NET was a provider network that served sites in the San Francisco Bay Area). In 1990, many providers served a limited geographic region and were thus known as regional networks. The regional networks were, in turn, connected by a nationwide backbone. In 1990, this backbone was funded by the National Science Foundation (NSF) and was therefore called the NSFNET backbone. Although the detail is not shown in this figure, the provider networks are typically built from a large number of point-to-point links (e.g., DS3 or OC-3 links) that connect to routers; similarly, each end user site is typically not a single network, but instead consists of multiple physical networks connected by routers and bridges. Notice in Figure 4.24 that each provider and end user are likely to be an ad- ministratively independent entity. This has some significant consequences on routing. For example, it is quite likely that different providers will have different ideas about 4.3 Global Internet 301 the best routing protocol to use within their network, and on how metrics should be assigned to links in their network. Because of this independence, each provider’s network is usually a single autonomous system (AS). We will define this term more precisely in Section 4.3.3, but for now it is adequate to think of an AS as a network that is administered independently of other ASs. The fact that the Internet has a discernible structure can be used to our advantage as we tackle the problem of scalability. In fact, we need to deal with two related scaling issues. The first is the scalability of routing. We need to find ways to minimize the number of network numbers that get carried around in routing protocols and stored in the routing tables of routers. The second is address utilization—that is, making sure that the IP address space does not get consumed too quickly. Throughout this section, we will see the principle of hierarchy used again and again to improve scalability. We begin with subnetting, which primarily deals with address space utilization. Next we introduce classless routing or supernetting, which tackles both address utilization and routing scalability. We then look at how hier- archy can be used to improve the scalability of routing, both through interdomain routing and within a single domain. Our final subsection looks at the emerging stan- dards for IP version 6, the invention of which was largely the result of scalability concerns. 4.3.1 Subnetting The original intent of IP addresses was that the network part would uniquely identify exactly one physical network. It turns out that this approach has a couple of draw- backs. Imagine a large campus that has lots of internal networks and that decides to connect to the Internet. For every network, no matter how small, the site needs at least a class C network address. Even worse, for any network with more than 255 hosts, they need a class B address. This may not seem like a big deal, and indeed it wasn’t when the Internet was first envisioned, but there are only a finite number of network numbers, and there are far fewer class B addresses than class Cs. Class B addresses tend to be in particularly high demand because you never know if your network might expand beyond 255 nodes, so it is easier to use a class B address from the start than to have to renumber every host when you run out of room on a class C network. The problem we observe here is address assignment inefficiency: A network with two nodes uses an entire class C network address, thereby wasting 253 perfectly useful addresses; a class B network with slightly more than 255 hosts wastes over 64,000 addresses. Assigning one network number per physical network, therefore, uses up the IP address space potentially much faster than we would like. While we would need to connect over 4 billion hosts to use up all the valid addresses, we only need to 302 4 Internetworking connect 214 (about 16,000) class B networks before that part of the address space runs out. Therefore, we would like to find some way to use the network numbers more efficiently. Assigning many network numbers has another drawback that becomes apparent when you think about routing. Recall that the amount of state that is stored in a node participating in a routing protocol is proportional to the number of other nodes, and that routing in an internet consists of building up forwarding tables that tell a router how to reach different networks. Thus, the more network numbers there are in use, the bigger the forwarding tables get. Big forwarding tables add cost to routers, and they are potentially slower to search than smaller tables for a given technology, so they degrade router performance. This provides another motivation for assigning network numbers carefully. Subnetting provides an elegantly simple way to reduce the total number of net- work numbers that are assigned. The idea is to take a single IP network number and allocate the IP addresses with that network number to several physical networks, which are now referred to as subnets. Several things need to be done to make this work. First, the subnets should be close to each other. This is because at a distant point in the Internet, they will all look like a single network, having only one network number between them. This means that a router will only be able to select one route to reach any of the subnets, so they had better all be in the same general direction. A perfect situation in which to use subnetting is a large campus or corporation that has many physical networks. From outside the campus, all you need to know to reach any subnet inside the campus is where the campus connects to the rest of the Internet. This is often at a single point, so one entry in your forwarding table will suffice. Even if there are multiple points at which the campus is connected to the rest of the Internet, knowing how to get to one point in the campus network is still a good start. The mechanism by which a single network number can be shared among multiple networks involves configuring all the nodes on each subnet with a subnet mask. With simple IP addresses, all hosts on the same network must have the same network number. The subnet mask enables us to introduce a subnet number; all hosts on the same physical network will have the same subnet number, which means that hosts may be on different physical networks but share a single network number. What the subnet mask effectively does is introduce another level of hierarchy into the IP address. For example, suppose that we want to share a single class B address among several physical networks. We could use a subnet mask of 255.255.255.0. (Subnet masks are written down just like IP addresses; this mask is therefore all 1s in the upper 24 bits and 0s in the lower 8 bits.) In effect, this means that the top 24 bits (where the mask has 1s) are now defined to be the network number, and the lower 8 bits (where the mask has 0s) are the host number. Since the top 16 bits identify the 4.3 Global Internet 303 Network number Host number Class B address Subnet mask (255.255.255.0) Subnetted address 111111111111111111111111 00000000 Network number Host IDSubnet ID Figure 4.25 Subnet addressing. Subnet mask: 255.255.255.128 Subnet number: 128.96.34.0 128.96.34.15 128.96.34.1 H1 R1 128.96.34.130 Subnet mask: 255.255.255.128 Subnet number: 128.96.34.128 128.96.34.129 128.96.34.139 R2 H2 128.96.33.1 128.96.33.14 Subnet mask: 255.255.255.0 Subnet number: 128.96.33.0 H3 Figure 4.26 An example of subnetting. network in a class B address, we may now think of the address as having not two parts but three: a network part, a subnet part, and a host part. That is, we have divided what used to be the host part into a subnet part and a host part. This is shown in Figure 4.25. What subnetting means to a host is that it is now configured with both an IP address and a subnet mask for the subnet to which it is attached. For example, host H1 in Figure 4.26 is configured with an address of 128.96.34.15 and a subnet mask of 255.255.255.128. (All hosts on a given subnet are configured with the same mask; 304 4 Internetworking i.e., there is exactly one subnet mask per subnet.) The bitwise AND of these two numbers defines the subnet number of the host and of all other hosts on the same subnet. In this case, 128.96.34.15 AND 255.255.255.128 equals 128.96.34.0, so this is the subnet number for the topmost subnet in the figure. When the host wants to send a packet to a certain IP address, the first thing it does is to perform a bitwise AND between its own subnet mask and the destination IP address. If the result equals the subnet number of the sending host, then it knows that the destination host is on the same subnet and the packet can be delivered directly over the subnet. If the results are not equal, the packet needs to be sent to a router to be forwarded to another subnet. For example, if H1 is sending to H2, then H1 ANDs its subnet mask (255.255.255.128) with the address for H2 (128.96.34.139) to obtain 128.96.34.128. This does not match the subnet number for H1 (128.96.34.0) so H1 knows that H2 is on a different subnet. Since H1 cannot deliver the packet to H2 directly over the subnet, it sends the packet to its default router R1. Note that ARP is largely unaffected by the change in address structure. Once a host or router figures out which node it needs to deliver a packet to on one of the networks to which it is attached, it performs ARP to find the MAC address for that node if necessary. The job of a router also changes when we introduce subnetting. Recall that, for simple IP, a router has a forwarding table that consists of entries of the form NetworkNum, NextHop. To support subnetting, the table must now hold entries of the form SubnetNumber, SubnetMask, NextHop. To find the right entry in the table, the router ANDs the packet’s destination address with the SubnetMask for each entry in turn; if the result matches the SubnetNumber of the entry, then this is the right entry to use, and it forwards the packet to the next hop router indicated. In the example network of Figure 4.26, router R1 would have the entries shown in Table 4.10. Continuing with the example of a datagram from H1 being sent to H2, R1 would AND H2’s address (128.96.34.139) with the subnet mask of the first entry (255.255.255.128) and compare the result (128.96.34.128) with the network number SubnetNumber SubnetMask NextHop 128.96.34.0 255.255.255.128 Interface 0 128.96.34.128 255.255.255.128 Interface 1 128.96.33.0 255.255.255.0 R2 Table 4.10 Example forwarding table with subnetting for Figure 4.26. 4.3 Global Internet 305 for that entry (128.96.34.0). Since this is not a match, it proceeds to the next entry. This time a match does occur, so R1 delivers the datagram to H2 using interface 1, which is the interface connected to the same network as H2. We can now describe the datagram forwarding algorithm in the following way: D = destination IP address for each forwarding table entry SubnetNumber, SubnetMask, NextHopD1 = SubnetMask & D if D1 = SubnetNumber if NextHop is an interface deliver datagram directly to destination else deliver datagram to NextHop (a router) Although not shown in this example, a default router would usually be included in the table and would be used if no explicit matches were found. We note in passing that a naive implementation of this algorithm—one involving repeated ANDing of the destination address with a subnet mask that may not be different every time, and a linear table search—would be very inefficient. A few fine points about subnetting need to be mentioned. We have already seen that the subnet mask does not need to align with a byte boundary, with the example mask of 255.255.255.128 (25 1s followed by 7 0s) used above. More confusingly, it is not even necessary for all the 1s in a subnet mask to be contiguous. For example, it would be quite possible to use a subnet mask of 255.255.1.0. All of the mechanisms described above should continue to work, but now you can’t look at a contiguous part of the IP address and say, “That is the subnet number.” This makes administration more difficult. It may also fail to work with implementations that assume that no one would use noncontiguous masks, and so it is not recommended in practice. We can also put multiple subnets on a single physical network. The effect of this would be to force hosts on the same network to talk to each other through a router, which might be useful for administrative purposes; for example, to provide isolation between different departments sharing a LAN. A third point to which we have alluded is that different parts of the internet see the world differently. From outside our hypothetical campus, routers see a single network. In the example above, routers outside the campus see the collection of networks in Figure 4.26 as just the network 128.96, and they keep one entry in their forwarding tables to tell them how to reach it. Routers within the campus, however, need to be able to route packets to the right subnet. Thus, not all parts of the internet see exactly the same routing information. The next section takes a closer look at how the propagation of routing information is done in the Internet. 306 4 Internetworking ◮ The bottom line is that subnetting helps solve our scalability problems in two ways. First, it improves our address assignment efficiency by letting us not use up an entire class C or class B address every time we add a new physical network. Second, it helps us aggregate information. From a reasonable distance, a complex collection of physical networks can be made to look like a single network, so that the amount of information that routers need to store to deliver datagrams to those networks can be reduced. 4.3.2 Classless Routing (CIDR) Classless interdomain routing (CIDR, pronounced “cider”) is a technique that addresses two scaling concerns in the Internet: the growth of backbone routing ta- bles as more and more network numbers need to be stored in them, and the potential for the 32-bit IP address space to be exhausted well before the four-billionth host is attached to the Internet. We have already mentioned the problem that would cause this address space exhaustion: address assignment inefficiency. The inefficiency arises because the IP address structure, with class A, B, and C addresses, forces us to hand out network address space in fixed-sized chunks of three very different sizes. A net- work with two hosts needs a class C address, giving an address assignment efficiency of 2/255 = 0.78%; a network with 256 hosts needs a class B address, for an effi- ciency of only 256/65,535 = 0.39%. Even though subnetting can help us to assign addresses carefully, it does not get around the fact that any autonomous system with more than 255 hosts, or an expectation of eventually having that many, wants a class B address. As it turns out, exhaustion of the IP address space centers on exhaustion of the class B network numbers. One way to deal with that would seem to be saying no to any AS that requests a class B address unless they can show a need for something close to 64K addresses, and instead giving them an appropriate number of class C addresses to cover the expected number of hosts. Since we would now be handing out address space in chunks of 256 addresses at a time, we could more accurately match the amount of address space consumed to the size of the AS. For any AS with at least 256 hosts (which means the majority of ASs), we can guarantee an address utilization of at least 50%, and typically much more. This solution, however, raises a problem that is at least as serious: excessive storage requirements at the routers. If a single AS has, say, 16 class C network numbers assigned to it, that means every Internet backbone router needs 16 entries in its routing tables for that AS. This is true even if the path to every one of those networks is the same. If we had assigned a class B address to the AS, the same routing information could be stored in one table entry. However, our address assignment efficiency would then be only 16 × 255/65,536 = 6.2%. 4.3 Global Internet 307 CIDR, therefore, tries to balance the desire to minimize the number of routes that a router needs to know against the need to hand out addresses efficiently. To do this, CIDR helps us to aggregate routes. That is, it lets us use a single entry in a forwarding table to tell us how to reach a lot of different networks. As you may have guessed from the name, it does this by breaking the rigid boundaries between address classes. To understand how this works, consider our hypothetical AS with 16 class C network numbers. Instead of handing out 16 addresses at random, we can hand out a block of contiguous class C addresses. Suppose we assign the class C network numbers from 192.4.16 through 192.4.31. Observe that the top 20 bits of all the addresses in this range are the same (11000000 00000100 0001). Thus, what we have effectively created is a 20-bit network number—something that is between a class B network number and a class C number in terms of the number of hosts that it can support. In other words, we get both the high address efficiency of handing out addresses in chunks smaller than a class B network and a single network prefix that can be used in forwarding tables. Observe that for this scheme to work, we need to hand out blocks of class C addresses that share a common prefix, which means that each block must contain a number of class C networks that is a power of two. All we need now to make CIDR solve our problems is a routing protocol that can deal with these “classless” addresses, which means that it must understand that a network number may be of any length. Modern routing protocols (such as BGP-4, described below) do exactly that. The network numbers that are carried in such a routing protocol are represented simply by length, value pairs, where the length gives the number of bits in the network prefix—20 in the above example. Note that representing a network address in this way is similar to the mask, value approach used in subnetting, as long as masks consist of contiguous bits starting from the most significant bit. Also note that we used subnetting to share one address among multiple physical networks, while CIDR aims to collapse the multiple addresses that would be assigned to a single AS onto one address. The similarity between the two approaches is reflected in the original name for CIDR—supernetting. In fact, the ability to aggregate routes in the way that we have just shown is only the first step. Imagine an Internet service provider network, whose primary job is to provide Internet connectivity to a large number of corporations and campuses. If we assign network numbers to the corporations in such a way that all the different corporations connected to the provider network share a common address prefix, then we can get even greater aggregation of routes. Consider the example in Figure 4.27. The two corporations served by the provider network have been assigned adjacent 20- bit network prefixes. Since both of the corporations are reachable through the same provider network, it can advertise a single route to both of them by just advertising the common 19-bit prefix they share. In general, it is possible to aggregate routes repeatedly if addresses are assigned carefully. This means that we need to pay attention to which 308 4 Internetworking Border gateway (advertises path to 11000000000001) Regional network Corporation X (11000000000001000001) Corporation Y (11000000000001000000) Figure 4.27 Route aggregation with CIDR. provider a corporation is attached to before assigning it an address if this scheme is to work. One way to accomplish that is to assign a portion of address space to the provider and then to let the network provider assign addresses from that space to its customers. IP Forwarding Revisited In all our discussion of IP forwarding so far, we have assumed that we could find the network number in a packet and then look up that number in a forwarding table. However, now that we have introduced CIDR, we need to reexamine this assumption. CIDR means that prefixes may be of any length, from 2 to 32 bits. Furthermore, it is sometimes possible to have prefixes in the forwarding table that “overlap,” in the sense that some addresses may match more than one prefix. For example, we might find both 171.69 (a 16-bit prefix) and 171.69.10 (a 24-bit prefix) in the forwarding table of a single router. In this case, a packet destined to, say, 171.69.10.5 clearly matches both prefixes. The rule in this case is based on the principle of “longest match”; that is, the packet matches the longest prefix, which would be 171.69.10 in this example. On the other hand, a packet destined to 171.69.20.5 would match 171.69 and not 171.69.10, and in the absence of any other matching entry in the routing table, 171.69 would be the longest match. The task of efficiently finding the longest match between an IP address and the variable-length prefixes in a forwarding table has been a fruitful field of research in recent years, and the “Further Reading” section of this chapter provides some references. The most well-known algorithm uses an approach known as a PATRICIA tree, which was actually developed well in advance of CIDR. 4.3.3 Interdomain Routing (BGP) At the beginning of this section we introduced the notion that the Internet is organized as autonomous systems, each of which is under the control of a single administrative entity. A corporation’s complex internal network might be a single AS, as may the 4.3 Global Internet 309 R1 Autonomous system 1 R2 R3 Autonomous system 2 R4 R5 R6 Figure 4.28 A network with two autonomous systems. network of a single Internet service provider. Figure 4.28 shows a simple network with two autonomous systems. The basic idea behind autonomous systems is to provide an additional way to hierarchically aggregate routing information in a large internet, thus improving scal- ability. We now divide the routing problem into two parts: routing within a single autonomous system and routing between autonomous systems. Since another name for autonomous systems in the Internet is routing domains, we refer to the two parts of the routing problem as interdomain routing and intradomain routing. In addition to improving scalability, the AS model decouples the intradomain routing that takes place in one AS from that taking place in another. Thus, each AS can run whatever intradomain routing protocols it chooses. It can even use static routes or multiple pro- tocols if desired. The interdomain routing problem is then one of having different ASs share reachability information with each other. One feature of the autonomous system idea is that it enables some ASs to dra- matically reduce the amount of routing information they need to care about by using default routes. For example, if a corporate network is connected to the rest of the 310 4 Internetworking Internet by a single router (this router is typically called a border router since it sits at the boundary between the AS and the rest of the Internet), then it is pretty easy for a host or router inside the autonomous system to figure out where it should send packets that are headed for a destination outside of this AS—they first go to the AS’s border router. This is the default route. Similarly, a regional Internet service provider can keep track of how to reach the networks of all its directly connected customers and can have a default route to some other provider (typically a backbone provider) for everyone else. Of course, this passing of the buck has to stop at some point; eventually the packet should reach a router connected to a backbone network that knows how to reach everything. Managing the amount of routing information in the backbones is an important issue that we discuss below. There have been two major interdomain routing protocols in the recent history of the Internet. The first was the Exterior Gateway Protocol (EGP). EGP had a number of limitations, perhaps the most severe of which was that it constrained the topology of the Internet rather significantly. EGP basically forced a treelike topology onto the Internet, or to be more precise, it was designed when the Internet had a treelike topology, such as that illustrated in Figure 4.24. EGP did not allow for the topology to become more general. Note that in this simple treelike structure, there is a single backbone, and autonomous systems are connected only as parents and children and not as peers. The replacement for EGP is the Border Gateway Protocol (BGP), which is in its fourth version at the time of this writing (BGP-4). BGP is also known for being rather complex. This section presents the highlights of BGP-4. As a starting position, BGP assumes that the Internet is an arbitrarily intercon- nected set of ASs. This model is clearly general enough to accommodate non-tree- structured internetworks, like the simplified picture of today’s multibackbone Internet shown in Figure 4.29.6 Unlike the simple tree-structured Internet shown in Figure 4.24, today’s Internet consists of an interconnection of multiple backbone networks (they are usually called service provider networks, and they are operated by private companies rather than the government), and sites are connected to each other in arbitrary ways. Some large corporations connect directly to one or more of the backbones, while others connect to smaller, nonbackbone service providers. Many service providers exist mainly to provide service to “consumers” (i.e., individuals with PCs in their homes), and these providers must also connect to the backbone providers. Often many providers arrange to interconnect with each other at a single “peering point.” In short, it is hard to discern much structure at all in today’s Internet. 6In an interesting stretch of metaphor, the Internet now has multiple backbones, having had only one for most of its early life. The authors know of no other animal that has this characteristic. 4.3 Global Internet 311 Backbone service provider Peering point Peering point Large corporation Large corporation Small corporation “Consumer” ISP “Consumer” ISP “Consumer” ISP Figure 4.29 Today’s multibackbone Internet. Given this rough sketch of the Internet, if we define local traffic as traffic that originates at or terminates on nodes within an AS, and transit traffic as traffic that passes through an AS, we can classify ASs into three types: ■ Stub AS: an AS that has only a single connection to one other AS; such an AS will only carry local traffic. The small corporation in Figure 4.29 is an example of a stub AS. ■ Multihomed AS: an AS that has connections to more than one other AS but that refuses to carry transit traffic; for example, the large corporation at the top of Figure 4.29. ■ Transit AS: an AS that has connections to more than one other AS and that is designed to carry both transit and local traffic, such as the backbone providers in Figure 4.29. Whereas the discussion of routing in Section 4.2 focused on finding optimal paths based on minimizing some sort of link metric, the problem of interdomain routing turns out to be so difficult that the goals are more modest. First and foremost, the goal is to find any path to the intended destination that is loop-free. That is, we are more concerned with reachability than optimality. Finding a path that is anywhere close to optimal is considered a great achievement. We will see why this is so as we look at the details of BGP. There are a few reasons why interdomain routing is hard. The first is simply a matter of scale. An Internet backbone router must be able to forward any packet destined anywhere in the Internet. That means having a routing table that will provide a match for any valid IP address. While CIDR has helped to control the number of 312 4 Internetworking distinct prefixes that are carried in the Internet’s backbone routing, there is inevitably a lot of routing information to pass around—on the order of 140,000 prefixes at the time of writing. The second challenge in interdomain routing arises from the autonomous nature of the domains. Note that each domain may run its own interior routing protocols and use any scheme it chooses to assign metrics to paths. This means that it is impossible to calculate meaningful path costs for a path that crosses multiple ASs. A cost of 1000 across one provider might imply a great path, but it might mean an unacceptably bad one from another provider. As a result, interdomain routing advertises only “reacha- bility.” The concept of reachability is basically a statement that “you can reach this network through this AS.” This means that for interdomain routing to pick an optimal path is essentially impossible. The third challenge involves the issue of trust. Provider A might be unwilling to believe certain advertisements from provider B for fear that provider B will advertise erroneous routing information. For example, trusting provider B when he advertises a great route to anywhere in the Internet can be a disastrous choice if provider B turns out to have made a mistake configuring his routers or to have insufficient capacity to carry the traffic. Closely related to this issue is the need to support very flexible policies in inter- domain routing. One common policy is the prevention of transit traffic. For example, the multihomed corporation in Figure 4.29 may not wish to carry any traffic between the two providers to whom it connects. As a more complex example, provider A might wish to implement policies that say, “Use provider B only to reach these addresses,” “Use the path that crosses the fewest number of ASs,” or “Use AS x in preference to AS y.” The goal is to specify policies that lead to “good” paths, if not to optimal ones. When configuring BGP, the administrator of each AS picks at least one node to be a “BGP speaker,” which is essentially a spokesperson for the entire AS. That BGP speaker establishes BGP sessions to other BGP speakers in other ASs. These sessions are used to exchange reachability information among ASs. In addition to the BGP speakers, the AS has one or more border “gateways,” which need not be the same as the speakers. The border gateways are the routers through which packets enter and leave the AS. In our simple example in Figure 4.28, routers R2 and R4 would be border gateways. Note that we have avoided using the word “gateway” until this point because it tends to be confusing. We can’t avoid it here, given the name of the protocol we are describing. The important point to understand here is that, in the context of interdomain routing, a border gateway is simply an IP router that is charged with the task of forwarding packets between ASs. BGP does not belong to either of the two main classes of routing protocols (distance-vector and link-state protocols) described in Section 4.2. Unlike these 4.3 Global Internet 313 Regional provider A (AS 2) Regional provider B (AS 3) Customer P (AS 4) Customer Q (AS 5) Customer R (AS 6) Customer S (AS 7) 128.96 192.4.153 192.4.32 192.4.3 192.12.69 192.4.54 192.4.23 Backbone network (AS 1) Figure 4.30 Example of a network running BGP. protocols, BGP advertises complete paths as an enumerated list of ASs to reach a particular network. This is necessary to enable the sorts of policy decisions described above to be made in accordance with the wishes of a particular AS. It also enables routing loops to be readily detected. To see how this works, consider the example network in Figure 4.30. Assume that the providers are transit networks, while the customer networks are stubs. A BGP speaker for the AS of provider A (AS 2) would be able to advertise reachability in- formation for each of the network numbers assigned to customers P and Q. Thus, it would say, in effect, “The networks 128.96, 192.4.153, 192.4.32, and 192.4.3 can be reached directly from AS 2.” The backbone network, on receiving this advertise- ment, can advertise, “The networks 128.96, 192.4.153, 192.4.32, and 192.4.3 can be reached along the path AS 1, AS 2.” Similarly, it could advertise, “The networks 192.12.69, 192.4.54, and 192.4.23 can be reached along the path AS 1, AS 3.” An important job of BGP is to prevent the establishment of looping paths. For example, consider three interconnected ASs, 1, 2, and 3. Suppose AS 1 learns that it can reach network 10.0.1 through AS 2, so it advertises this fact to AS 3, who in turn advertises it back to AS 2. AS 2 could now decide that AS 3 was the place to send packets destined for 10.0.1; AS 3 sends them to AS 1; AS 1 sends them back to AS 2; and they would loop forever. This is prevented by carrying the complete AS path in the routing messages. In this case, the advertisement received by AS 2 from AS 3 would contain an AS path of AS 3, AS 1, AS 2. AS 2 sees itself in this path, and thus concludes that this is not a useful path for it to use. It should be apparent that the AS numbers carried in BGP need to be unique. For example, AS 2 can only recognize itself in the AS path in the above example if 314 4 Internetworking no other AS identifies itself in the same way. AS numbers are 16-bit numbers assigned by a central authority to assure uniqueness. While 16 bits only allows about 65,000 ASs, which might not seem like a lot, we note that stub ASs do not need a unique AS number, and this covers the overwhelming majority of nonprovider networks. We should note that a given AS will only advertise routes that it considers good enough for itself. That is, if a BGP speaker has a choice of several different routes to a destination, it will choose the best one according to its own local policies, and then that will be the route it advertises. Furthermore, a BGP speaker is under no obligation to advertise any route to a destination, even if it has one. This is how an AS can implement a policy of not providing transit—by refusing to advertise routes to prefixes that are not contained within that AS, even if it knows how to reach them. In addition to advertising paths, BGP speakers need to be able to cancel previously advertised paths if a critical link or node on a path goes down. This is done with a form of negative advertisement known as a withdrawn route. Both positive and negative reachability information are carried in a BGP update message, the format of which is shown in Figure 4.31. (Note that the fields in this figure are multiples of 16 bits, unlike other packet formats in this chapter.) One point to note about BGP-4 is that it was designed to cope with the classless addresses described in Section 4.3.2. This means that the “networks” that are adver- tised in BGP are actually prefixes of any length. Thus, the updates contain both the prefix itself and its length in bits. When writing these down, it is common to write prefix/length. For example, a CIDR prefix that begins 192.4.16 and is 20 bits long would be written as 192.4.16/20. 015 Unfeasible routes length Withdrawn routes (variable) Total path attribute length Path attributes (variable) Network layer reachability info (variable) Figure 4.31 BGP-4 update packet format. 4.3 Global Internet 315 A final point to note is that BGP is defined to run on top of TCP, the reliable transport protocol described in Section 5.2. Because BGP speakers can count on TCP to be reliable, this means that any information that has been sent from one speaker to another does not need to be sent again. Thus, as long as nothing has changed, a BGP speaker can simply send an occasional “keepalive” message that says, in effect, “I’m still here and nothing has changed.” If that router were to crash, it would stop sending the keepalives, and the other routers that had learned routes from it would know that those routes were no longer valid. We will not delve further into the details of BGP-4, except to point out that all the protocol does is specify how reachability information should be exchanged among autonomous systems. BGP speakers obtain enough information by this exchange to calculate loop-free routes to all reachable networks, but how they choose the “best” routes is largely left to the policies of the AS.◮ Let’s return to the real question: How does all this help us to build scalable networks? First, the number of nodes participating in BGP is on the order of the number of ASs, which is much smaller than the number of networks. Second, finding a good interdomain route is only a matter of finding a path to the right border router, of which there are only a few per AS. Thus, we have neatly subdivided the routing problem into manageable parts, once again using a new level of hierarchy to increase scalability. The complexity of interdomain routing is now on the order of the number of ASs, and the complexity of intradomain routing is on the order of the number of networks in a single AS. Integrating Interdomain and Intradomain Routing While the preceding discussion illustrates how a BGP speaker learns interdomain rout- ing information, the question still remains as to how all the other routers in a domain get this information. There are several ways this problem can be addressed. We have already alluded to one very simple situation, which is also very common. In the case of a stub AS that only connects to other ASs at a single point, the border router is clearly the only choice for all routes that are outside the AS. Such a router can “inject” a default route into the intradomain routing protocol. In effect, this is a statement that any network that has not been explicitly advertised in the intradomain protocol is reachable through the border router. Recall from the discussion of IP for- warding in Section 4.1 that the default entry in the forwarding table comes after all the more specific entries, and it matches anything that failed to match a specific entry. The next step up in complexity is to have the border routers inject specific routes they have learned from outside the AS. Consider, for example, the border router of a provider AS that connects to a customer AS. That router could learn that the network prefix 192.4.54/24 is located inside the customer AS, either through BGP or because 316 4 Internetworking the information is configured into the border router. It could inject a route to that prefix into the routing protocol running inside the provider AS. This would be an advertisement of the sort “I have a link to 192.4.54/24 of cost X.” This would cause other routers in the provider AS to learn that this border router is the place to send packets destined for that prefix. The final level of complexity comes in backbone networks, which learn so much routing information from BGP that it becomes too costly to inject it into the intrado- main protocol. For example, if a border router wants to inject 10,000 prefixes that it learned about from another AS, it will have to send very big link-state packets to the other routers in that AS, and their shortest-path calculations are going to become very complex. For this reason, the routers in a backbone network use a variant of BGP called interior BGP (IBGP) to effectively redistribute the information that is learned by the BGP speakers at the edges of the AS to all the other routers in the AS. IBGP enables any router in the AS to learn the best border router to use when sending a packet to any address. At the same time, each router in the AS keeps track of how to get to each border router using a conventional intradomain protocol with no injected information. By combining these two sets of information, each router in the AS is able to determine the appropriate next hop for all prefixes. 4.3.4 Routing Areas As if we didn’t already have enough hierarchy, link-state intradomain routing protocols provide a means to partition a routing domain into subdomains called areas. (The terminology varies somewhat among protocols; we use the OSPF terminology here.) By adding this extra level of hierarchy, we enable single domains to grow larger without overburdening the intradomain routing protocols. An area is a set of routers that are administratively configured to exchange link- state information with each other. There is one special area—the backbone area, also known as area 0. An example of a routing domain divided into areas is shown in Figure 4.32. Routers R1, R2, and R3 are members of the backbone area. They are also members of at least one nonbackbone area; R1 is actually a member of both area 1 and area 2. A router that is a member of both the backbone area and a nonbackbone area is an area border router (ABR). Note that these are distinct from the routers that are at the edge of an AS, which are referred to as AS border routers for clarity. Routing within a single area is exactly as described in Section 4.2.3. All the routers in the area send link-state advertisements to each other, and thus develop a complete, consistent map of the area. However, the link-state advertisements of routers that are not area border routers do not leave the area in which they originated. This has the effect of making the flooding and route calculation processes considerably more scalable. For example, router R4 in area 3 will never see a link-state advertisement 4.3 Global Internet 317 Area 1 Area 0 Area 3 Area 2 R9 R8 R7 R1 R5R6 R4 R3 R2 Figure 4.32 A domain divided into areas. from router R8 in area 1. As a consequence, it will know nothing about the detailed topology of areas other than its own. How, then, does a router in one area determine the right next hop for a packet destined to a network in another area? The answer to this becomes clear if we imagine the path of a packet that has to travel from one nonbackbone area to another as being split into three parts. First, it travels from its source network to the backbone area, then it crosses the backbone, then it travels from backbone to destination network. To make this work, the area border routers summarize routing information that they have learned from one area and make it available in their advertisements to other areas. For example, R1 receives link-state advertisements from all the routers in area 1 and can thus determine the cost of reaching any network in area 1. When R1 sends link-state advertisements into area 0, it advertises the costs of reaching the networks in area 1 much as if all those networks were directly connected to R1. This enables all the area 0 routers to learn the cost to reach all networks in area 1. The area border routers then summarize this information and advertise it into the nonbackbone areas. Thus, all routers learn how to reach all networks in the domain. Note that in the case of area 2, there are two ABRs, and that routers in area 2 will thus have to make a choice as to which one they use to reach the backbone. This is easy enough, since both R1 and R2 will be advertising costs to various networks, so that it will become clear which is the better choice as the routers in area 2 run their shortest-path algorithm. For example, it is pretty clear that R1 is going to be a better choice than R2 for destinations in area 1. 318 4 Internetworking When dividing a domain into areas, the network administrator makes a trade- off between scalability and optimality of routing. The use of areas forces all packets traveling from one area to another to go via the backbone area, even if a shorter path might have been available. For example, even if R4 and R5 were directly connected, packets would not flow between them because they are in different nonbackbone areas. It turns out that the need for scalability is often more important than the need to use the absolute shortest path.◮ This illustrates an important principle in network design. There is frequently a trade-off between some sort of optimality and scalability. When hierarchy is in- troduced, information is hidden from some nodes in the network, hindering their ability to make perfectly optimal decisions. However, information hiding is essential to scalability, since it saves all nodes from having global knowledge. It is invariably true in large networks that scalability is a more pressing design goal than perfect optimality. Finally, we note that there is a trick by which network administrators can more flexibly decide which routers go in area 0. This trick uses the idea of a “virtual link” between routers. Such a virtual link is obtained by configuring a router that is not directly connected to area 0 to exchange backbone routing information with a router that is. For example, a virtual link could be configured from R8 to R1, thus making R8 part of the backbone. R8 would now participate in link-state advertisement flooding with the other routers in area 0. The cost of the virtual link from R8 to R1 is determined by the exchange of routing information that takes place in area 1. This technique can help to improve the optimality of routing. 4.3.5 IP Version 6 (IPv6) In many respects, the motivation for a new version of IP is the same as the motivation for the techniques described in the last section: to deal with scaling problems caused by the Internet’s massive growth. Subnetting and CIDR have helped to contain the rate at which the Internet address space is being consumed (the address depletion problem) and have also helped to control the growth of routing table information needed in the Internet’s routers (the routing information problem). However, there will come a point at which these techniques are no longer adequate. In particular, it is virtually impossible to achieve 100% address utilization efficiency, so the address space will be exhausted well before the four-billionth host is connected to the Internet. Even if we were able to use all four billion addresses, it’s not too hard to imagine ways that number could be exhausted, such as the assignment of IP addresses to set-top boxes for cable TV or to electricity meters. All of these possibilities argue that a bigger address space than that provided by 32 bits will eventually be needed. 4.3 Global Internet 319 Historical Perspective The IETF began looking at the problem of expanding the IP address space in 1991, and several alternatives were proposed. Since the IP address is carried in the header of every IP packet, increasing the size of the address dictates a change in the packet header. This means a new version of the Internet Protocol, and as a consequence, a need for new software for every host and router in the Internet. This is clearly not a trivial matter—it is a major change that needs to be thought about very carefully. The effort to define a new version of IP was known as IP Next Generation, or IPng. As the work progressed, an official IP version number was assigned, so IPng is now known as IPv6. Note that the version of IP discussed so far in this chapter is version 4 (IPv4). The apparent discontinuity in numbering is the result of version number 5 being used for an experimental protocol some years ago. The significance of the change to a new version of IP caused a snowball effect. The general feeling among network designers was that if you are going to make a change of this magnitude, you might as well fix as many other things in IP as possible at the same time. Consequently, the IETF solicited white papers from anyone who cared to write one, asking for input on the features that might be desired in a new version of IP. In addition to the need to accommodate scalable routing and addressing, some of the other wish list items for IPng were ■ support for real-time services ■ security support ■ autoconfiguration (i.e., the ability of hosts to automatically configure them- selves with such information as their own IP address and domain name) ■ enhanced routing functionality, including support for mobile hosts It is interesting to note that while many of these features were absent from IPv4 at the time IPv6 was being designed, support for all of them has made its way into IPv4 in recent years. In addition to the wish list, one absolutely nonnegotiable feature for IPng was that there must be a transition plan to move from the current version of IP (version 4) to the new version. With the Internet being so large and having no centralized control, it would be completely impossible to have a “flag day” on which everyone shut down their hosts and routers and installed a new version of IP. Thus, there will probably be a long transition period in which some hosts and routers will run IPv4 only, some will run IPv4 and IPv6, and some will run IPv6 only. The IETF appointed a committee called the IPng Directorate to collect all the inputs on IPng requirements and to evaluate proposals for a protocol to become IPng. Over the life of this committee there were a number of proposals, some of which 320 4 Internetworking merged with other proposals, and eventually one was chosen by the Directorate to be the basis for IPng. That proposal was called SIPP (Simple Internet Protocol Plus). SIPP originally called for a doubling of the IP address size to 64 bits. When the Directorate selected SIPP, they stipulated several changes, one of which was another doubling of the address to 128 bits (16 bytes). It was around this time that the version number 6 was assigned. The rest of this section describes some of the main features of IPv6. At the time of this writing, most of the key specifications for IPv6 are Proposed or Draft Standards in the IETF. Addresses and Routing First and foremost, IPv6 provides a 128-bit address space, as opposed to the 32 bits of version 4. Thus, while version 4 can potentially address four billion nodes if address assignment efficiency reaches 100%, IPv6 can address 3.4×1038 nodes, again assuming 100% efficiency. As we have seen, though, 100% efficiency in address assignment is not likely. Some analysis of other addressing schemes, such as those of the French and U.S. telephone networks, as well as that of IPv4, have turned up some empirical numbers for address assignment efficiency. Based on the most pessimistic estimates of efficiency drawn from this study, the IPv6 address space is predicted to provide over 1500 addresses per square foot of the earth’s surface, which certainly seems like it should serve us well even when toasters on Venus have IP addresses. Address Space Allocation IPv6 addresses do not have classes, but the address space is still subdivided in vari- ous ways based on the leading bits. Rather than specifying different address classes, the leading bits specify different uses of the IPv6 address. The current assignment of prefixes is listed in Table 4.11. This allocation of the address space turns out to be easier to explain than it looks. First, the entire functionality of IPv4’s three main address classes (A, B, and C) is contained inside the 001 prefix. Aggregatable Global Unicast Addresses, as we will see shortly, are a lot like classless IPv4 addresses, only much longer. These are the main ones of interest at this point, with one-eighth of the address space allocated to this important form of address. Obviously, large chunks of address space have been left unassigned to allow for future growth and new features. Two portions of the address space (0000 001 and 0000 010) have been reserved for encoding of other (non-IP) address schemes. NSAP addresses are used by the ISO protocols, and IPX addresses are used by Novell’s network-layer protocol. The idea behind “link local use” addresses is to enable a host to construct an address that will work on the network to which it is connected without being concerned about global uniqueness of the address. This may be useful for autoconfiguration, as 4.3 Global Internet 321 Prefix Use 0000 0000 Reserved 0000 0001 Unassigned 0000 001 Reserved for NSAP allocation 0000 010 Reserved for IPX allocation 0000 011 Unassigned 0000 1 Unassigned 0001 Unassigned 001 Aggregatable Global Unicast Addresses 010 Unassigned 011 Unassigned 100 Unassigned 101 Unassigned 110 Unassigned 1110 Unassigned 1111 0 Unassigned 1111 10 Unassigned 1111 110 Unassigned 1111 1110 0 Unassigned 1111 1110 10 Link local use addresses 1111 1110 11 Site local use addresses 1111 1111 Multicast addresses Table 4.11 Address prefix assignments for IPv6. 322 4 Internetworking we will see below. Similarly, the “site local use” addresses are intended to allow valid addresses to be constructed on a site (e.g., a private corporate network) that is not con- nected to the larger Internet; again, global uniqueness need not be an issue. Finally, the multicast address space is for multicast, thereby serving the same role as class D addresses in IPv4. Note that mul- ticast addresses are easy to distinguish— they start with a byte of all 1s. We will see how these addresses are used in Section 4.4. Within the reserved address space (ad- dresses beginning with a byte of 0s) are some important special types of addresses. A node may be assigned an “IPv4-compatible IPv6 address” by zero-extending a 32-bit IPv4 address to 128 bits. A node that is only capa- ble of understanding IPv4 can be assigned an “IPv4-mapped IPv6 address” by prefix- ing the 32-bit IPv4 address with 2 bytes of all 1s and then zero-extending the result to 128 bits. These two special address types have uses in the IPv4-to-IPv6 transition (see the sidebar on this topic). Address Notation Just as with IPv4, there is some special not- ation for writing down IPv6 addresses. The standard representation is x:x:x:x:x:x:x:x, where each “x” is a hexadecimal represen- tation of a 16-bit piece of the address. An example would be 47CD:1234:4422:ACO2:0022:1234:A456:0124 Any IPv6 address can be written using this notation. Since there are a few spe- cial types of IPv6 addresses, there are some special notations that may be helpful in Transition from IPv4 to IPv6 The most important idea behind the transition from IPv4 to IPv6 is that the Internet is far too big and decentralized to have a “flag day”—one specified day on which every host and router is upgraded from IPv4 to IPv6. Thus, IPv6 needs to be deployed incrementally in such a way that hosts and routers that only understand IPv4 can con- tinue to function for as long as pos- sible. Ideally, IPv4 nodes should be able to talk to other IPv4 nodes and some set of other IPv6-capable nodes indefinitely. Also, IPv6 hosts should be capable of talking to other IPv6 nodes even when some of the infrastructure between them may only support IPv4. Two major mechanisms have been defined to help this transition: dual-stack operation and tunneling. The idea of dual stacks is fairly straightforward: IPv6 nodes run both IPv6 and IPv4 and use the Version field to decide which stack should process an arriving packet. In this case, the IPv6 address could be unrelated to the IPv4 address, or it could be the “IPv4-mapped IPv6 address” described earlier in this section. The basic tunneling technique, in which an IP packet is sent as 4.3 Global Internet 323 the payload of another IP packet, was described in Section 4.1. For IPv6 transition, tunneling is used to send an IPv6 packet over a piece of the network that only under- stands IPv4. This means that the IPv6 packet is encapsulated within an IPv4 header that has the address of the tunnel endpoint in its header, is transmitted across the IPv4-only piece of network, and then is decap- sulated at the endpoint. The end- point could be either a router or a host; in either case, it must be IPv6-capable to be able to process the IPv6 packet after decapsula- tion. If the endpoint is a host with an IPv4-mapped IPv6 address, then tunneling can be done automati- cally, by extracting the IPv4 address from the IPv6 address and using it to form the IPv4 header. Other- wise, the tunnel must be configured manually. In this case, the encap- sulating node needs to know the IPv4 address of the other end of the tunnel, since it cannot be ex- tracted from the IPv6 header. From the perspective of IPv6, the other end of the tunnel looks like a reg- ular IPv6 node that is just one hop away, even though there may be many hops of IPv4 infrastruc- ture between the tunnel endpoints. certain circumstances. For example, an ad- dress with a large number of contiguous 0s can be written more compactly by omitting all the 0 fields. Thus 47CD:0000:0000:0000:0000:0000:A456:0124 could be written 47CD::A456:0124 Clearly, this form of shorthand can only be used for one set of contiguous 0s in an address to avoid ambiguity. Since there are two types of IPv6 addresses that contain an embedded IPv4 address, these have their own special no- tation that makes extraction of the IPv4 address easier. For example, the “IPv4- mapped IPv6 address” of a host whose IPv4 address was 128.96.33.81 could be written as ::FFFF:128.96.33.81 That is, the last 32 bits are written in IPv4 notation, rather than as a pair of hexadec- imal numbers separated by a colon. Note that the double colon at the front indicates the leading 0s. Aggregatable Global Unicast Addresses By far the most important thing that IPv6 must provide when it is deployed is plain old unicast addressing. It must do this in a way that supports the rapid rate of addition of new hosts to the Internet and that allows routing to be done in a scalable way as the number of physical networks in the Inter- net grows. Thus, at the heart of IPv6 is the 324 4 Internetworking unicast address allocation plan that determines how addresses beginning with the 001 prefix will be assigned to service providers, autonomous systems, networks, hosts, and routers. In fact, the address allocation plan that is proposed for IPv6 unicast addresses is extremely similar to that being deployed with CIDR in IPv4. To understand how it works and how it provides scalability, it is helpful to define some new terms. We may think of a nontransit AS (i.e., a stub or multihomed AS) as a subscriber, and we may think of a transit AS as a provider. Furthermore, we may subdivide providers into direct and indirect. The former are directly connected to subscribers. The latter pri- marily connect other providers, are not connected directly to subscribers, and are often known as backbone networks. With this set of definitions, we can see that the Internet is not just an arbitrarily interconnected set of ASs; it has some intrinsic hierarchy. The difficulty is in making use of this hierarchy without inventing mechanisms that fail when the hierarchy is not strictly observed, as happened with EGP. For example, the distinction between direct and indirect providers becomes blurred when a subscriber connects to a backbone or when a direct provider starts connecting to many other providers. As with CIDR, the goal of the IPv6 address allocation plan is to provide aggre- gation of routing information to reduce the burden on intradomain routers. Again, the key idea is to use an address prefix—a set of contiguous bits at the most significant end of the address—to aggregate reachability information to a large number of net- works and even to a large number of ASs. The main way to achieve this is to assign an address prefix to a direct provider and then for that direct provider to assign longer prefixes that begin with that prefix to its subscribers. This is exactly what we observed in Figure 4.27. Thus, a provider can advertise a single prefix for all of its subscribers. Of course, the drawback is that if a site decides to change providers, it will need to obtain a new address prefix and renumber all the nodes in the site. This could be a colossal undertaking, enough to dissuade most people from ever changing providers. For this reason, there is ongoing research on other addressing schemes, such as geographic addressing, in which a site’s address is a function of its location rather than the provider to which it attaches. At present, however, provider-based addressing is necessary to make routing work efficiently. Note that while IPv6 address assignment is essentially equivalent to the way address assignment has happened in IPv4 since the introduction of CIDR, IPv6 has the significant advantage of not having a large installed base of assigned addresses to fit into its plans. One question is whether it makes sense for hierarchical aggregation to take place at other levels in the hierarchy. For example, should all providers obtain their address prefixes from within a prefix allocated to the backbone to which they connect? Given that most providers connect to multiple backbones, this probably doesn’t make sense. 4.3 Global Internet 325 125–m–n–o–pponm3 SubscriberIDProviderIDRegistryID010 InterfaceIDSubnetID Figure 4.33 An IPv6 provider-based unicast address. Also, since the number of providers is much smaller than the number of sites, the benefits of aggregating at this level are much less. One place where aggregation may make sense is at the national or continental level. Continental boundaries form natural divisions in the Internet topology, and if all addresses in Europe, for example, had a common prefix, then a great deal of aggregation could be done, so that most routers in other continents would only need one routing table entry for all networks with the Europe prefix. Providers in Europe would all select their prefixes such that they began with the European prefix. Using this scheme, an IPv6 address might look like Figure 4.33. The RegistryID might be an identifier assigned to a European address registry, with different IDs assigned to other continents or countries. Note that prefixes would be of different lengths under this scenario. For example, a provider with few customers could have a longer prefix (and thus less total address space available) than one with many customers. One tricky situation could occur when a subscriber is connected to more than one provider. Which prefix should the subscriber use for his or her site? There is no perfect solution to the problem. For example, suppose a subscriber is connected to two providers X and Y. If the subscriber takes his prefix from X, then Y has to advertise a prefix that has no relationship to its other subscribers and that as a consequence cannot be aggregated. If the subscriber numbers part of his AS with the prefix of X and part with the prefix of Y, he runs the risk of having half his site become unreachable if the connection to one provider goes down. One solution that works fairly well if X and Y have a lot of subscribers in common is for them to have three prefixes between them: one for subscribers of X only, one for subscribers of Y only, and one for the sites that are subscribers of both X and Y. Packet Format Despite the fact that IPv6 extends IPv4 in several ways, its header format is actually simpler. This simplicity is due to a concerted effort to remove unnecessary functionality from the protocol. Figure 4.34 shows the result. (For comparison with IPv4, see the header format shown in Figure 4.3.) As with many headers, this one starts with a Version field, which is set to 6 for IPv6. The Version field is in the same place relative to the start of the header as IPv4’s Version field so that header-processing software can immediately decide which header 326 4 Internetworking Version TrafficClass FlowLabel PayloadLen NextHeader HopLimit SourceAddress DestinationAddress 0 4 12 16 24 31 Next header/data Figure 4.34 IPv6 packet header. format to look for. The TrafficClass and FlowLabel fields both relate to quality of service issues, as discussed in Section 6.5. The PayloadLen field gives the length of the packet, excluding the IPv6 header, measured in bytes. The NextHeader field cleverly replaces both the IP options and the Protocol field of IPv4. If options are required, then they are carried in one or more special headers following the IP header, and this is indicated by the value of the NextHeader field. If there are no special headers, the NextHeader field is the demux key identifying the higher-level protocol running over IP (e.g., TCP or UDP); that is, it serves the same purpose as the IPv4 Protocol field. Also, fragmentation is now handled as an optional header, which means that the fragmentation-related fields of IPv4 are not included in the IPv6 header. The HopLimit field is simply the TTL of IPv4, renamed to reflect the way it is actually used. Finally, the bulk of the header is taken up with the source and destination ad- dresses, each of which is 16 bytes (128 bits) long. Thus, the IPv6 header is always 40 bytes long. Considering that IPv6 addresses are four times longer than those of IPv4, this compares quite well with the IPv4 header, which is 20 bytes long in the absence of options. 4.3 Global Internet 327 The way that IPv6 handles options is quite an improvement over IPv4. In IPv4, if any options were present, every router had to parse the entire options field to see if any of the options were relevant. This is because the options were all buried at the end of the IP header, as an unordered collection of type, length, value tuples. In contrast, IPv6 treats options as extension headers that must, if present, appear in a specific order. This means that each router can quickly determine if any of the options are relevant to it; in most cases, they will not be. Usually, this can be determined by just looking at the NextHeader field. The end result is that option processing is much more efficient in IPv6, which is an important factor in router performance. In addition, the new formatting of options as extension headers means that they can be of arbitrary length, whereas in IPv4 they were limited to 44 bytes at most. We will see how some of the options are used below. Each option has its own type of extension header. The type of each extension header is identified by the value of the NextHeader field in the header that precedes it, and each extension header contains a NextHeader field to identify the header following it. The last extension header will be followed by a transport-layer header (e.g., TCP), and in this case the value of the NextHeader field is the same as the value of the Protocol field in an IPv4 header. Thus, the NextHeader field does double duty; either it may identify the type of extension header to follow, or, in the last extension header, it serves as a demux key to identify the higher-layer protocol running over IPv6. Consider the example of the fragmentation header, shown in Figure 4.35. This header provides functionality similar to the fragmentation fields in the IPv4 header de- scribed in Section 4.1.2, but it is only present if fragmentation is necessary. Assuming it is the only extension header present, then the NextHeader field of the IPv6 header would contain the value 44, which is the value assigned to indicate the fragmentation header. The NextHeader field of the fragmentation header itself contains a value describing the header that follows it. Again, assuming no other extension headers are present, then the next header might be the TCP header, which results in NextHeader containing the value 6, just as the Protocol field would in IPv4. If the fragmentation header were followed by, say, an authentication header, then the fragmentation header’s NextHeader field would contain the value 51. NextHeader Reserved Offset RES M Ident 0 8 16 29 31 Figure 4.35 IPv6 fragmentation extension header. 328 4 Internetworking Autoconfiguration While the Internet’s growth has been im- pressive, one factor that has inhibited faster acceptance of the technology is the fact that getting connected to the Internet has typi- cally required a fair amount of system ad- ministration expertise. In particular, every host that is connected to the Internet needs to be configured with a certain minimum amount of information, such as a valid IP address, a subnet mask for the link to which it attaches, and the address of a name server. Thus, it has not been possible to unpack a new computer and connect it to the Internet without some preconfiguration. One goal of IPv6, therefore, is to provide support for autoconfiguration, sometimes referred to as “plug-and-play” operation. As we saw in Section 4.1.6, auto- configuration is possible for IPv4, but it depends on the existence of a server that is configured to hand out addresses and other configuration information to DHCP clients. The longer address format in IPv6 helps provide a useful, new form of autoconfiguration called stateless auto- configuration, which does not require a server. Recall that IPv6 unicast addresses are hierarchical, and that the least significant portion is the interface ID. Thus, we can subdivide the autoconfiguration problem into two parts: 1 Obtain an interface ID that is unique on the link to which the host is at- tached. 2 Obtain the correct address prefix for this subnet. Network Address Translation While IPv6 was motivated by a concern that increased usage of IP would lead to exhaustion of the address space, another technology has become popular as a way to conserve IP address space. That technology is network address translation (NAT), and it is pos- sible that its widespread use will significantly delay the need to de- ploy IPv6. NAT is often viewed as “architecturally impure,” but it is also a fact of networking life that cannot be ignored. The basic idea behind NAT is that all the hosts that might communicate with each other over the Internet do not need to have globally unique addresses. Instead, a host could be assigned a “pri- vate address” that is not necessar- ily globally unique, but is unique within some more limited scope; for example, within the corporate net- work where the host resides. The class A network number 10 is often used for this purpose, since that network number was assigned to the ARPANET and is no longer in use as a globally unique address. As long as the host communicates only with other hosts in the cor- porate network, a locally unique address is sufficient. If it should want to communicate with a host 4.3 Global Internet 329 outside the corporate network, it does so via a “NAT box”—a de- vice that is able to translate from the private address used by the host to some globally unique address that is assigned to the NAT box. Since it’s likely that a small sub- set of the hosts in the corpora- tion need the services of the NAT box at any one time, the NAT box might be able to get by with a small pool of globally unique addresses, much smaller than the number of addresses that would be needed if every host in the corporation had a globally unique address. So, we can imagine a NAT box receiving IP packets from a host inside the corporation and trans- lating the IP source address from some private address (say, 10.0.1.5) to a globally unique address (say, 171.69.210.246). When packets come back from the remote host addressed to 171.69. 210.246, the NAT box translates the destina- tion address to 10.0.1.5 and for- wards the packet on toward the host. The chief drawback of NAT is that it breaks a key assump- tion of the IP service model—that all nodes have globally unique ad- dresses. It turns out that lots of applications and protocols rely on The first part turns out to be rather easy, since every host on a link must have a unique link-level address. For example, all hosts on an Ethernet have a unique 48-bit Ethernet address. This can be turned into a valid link local use address by adding the appropriate prefix from Table 4.11 (1111 1110 10) followed by enough 0s to make up 128 bits. For some devices—for exam- ple, printers or hosts on a small router- less network that do not connect to any other networks—this address may be per- fectly adequate. Those devices that need a globally valid address depend on a router on the same link to periodically advertise the appropriate prefix for the link. Clearly, this requires that the router be configured with the correct address prefix, and that this prefix be chosen in such a way that there is enough space at the end (e.g., 48 bits) to attach an appropriate link-level address. The ability to embed link-level ad- dresses as long as 48 bits into IPv6 addresses was one of the reasons for choosing such a large address size. Not only does 128 bits allow the embedding, but it leaves plenty of space for the multilevel hierarchy of addressing that we discussed above. Advanced Routing Capabilities Another of IPv6’s extension headers is the routing header. In the absence of this header, routing for IPv6 differs very little from that of IPv4 under CIDR. The routing header contains a list of IPv6 addresses that rep- resent nodes or topological areas that the packet should visit en route to its des- tination. A topological area may be, for example, a backbone provider’s network. 330 4 Internetworking Specifying that packets must visit this net- work would be a way of implementing provider selection on a packet-by-packet basis. Thus, a host could say that it wants some packets to go through a provider that is cheap, others through a provider that provides high reliability, and still others through a provider that the host trusts to provide security. To provide the ability to specify topo- logical entities rather than individual nodes, IPv6 defines an anycast address. An anycast address is assigned to a set of interfaces, and packets sent to that address will go to the “nearest” of those interfaces, with nearest being determined by the routing protocols. For example, all the routers of a backbone provider could be assigned a single anycast address, which would be used in the routing header. The anycast address and the routing header are also expected to be used to pro- vide enhanced routing support to mobile hosts. The detailed mechanisms for provid- ing this support are still being defined. this assumption. In particular, many protocols that run over IP (e.g., application protocols) carry IP addresses in their mes- sages. These addresses also need to be translated by a NAT box if the higher-layer protocol is to work properly, and thus NAT boxes become much more com- plex than simple IP header trans- lators. They potentially need to un- derstand an ever-growing number of higher-layer protocols. This in turn presents an obstacle to deploy- ment of new applications. It is probably safe to say that networks would be better off without NAT, but its disappearance seems un- likely. Widespread deployment of IPv6 would almost certainly help. Other Features As mentioned at the beginning of this section, the primary motivation behind the devel- opment of IPv6 was to support the continued growth of the Internet. Once the IP header had to be changed for the sake of the addresses, however, the door was open for a wide variety of other changes, two of which we have just described—autoconfiguration and source-directed routing. IPv6 includes several additional features, most of which are covered elsewhere in this book—mobility is discussed in Section 4.2.5, network secu- rity is the topic of Chapter 8, and a new service model proposed for the Internet is described in Section 6.5. It is interesting to note that, in most of these areas, the IPv4 and IPv6 capabilities have become virtually indistinguishable, so that the main driver for IPv6 remains the need for larger addresses. 4.4 Multicast 331 4.4 Multicast As we saw in Chapter 2, multiaccess networks like Ethernet and token rings implement multicast in hardware. This section describes how to extend multicast, in software, across an internetwork of such networks. The approach described in this section is based on an implementation of multicast used in the current Internet (IPv4). Multicast will also be supported in the next generation of IP (IPv6), with the major differences being restricted to the address format. The motivation for developing multicast is that there are applications that want to send a packet to more than one destination host. Instead of forcing the source host to send a separate packet to each of the destination hosts, we want the source to be able to send a single packet to a multicast address, and for the network—or internet, in this case—to deliver a copy of that packet to each of a group of hosts. Hosts can then choose to join or leave this group at will, without synchronizing or negotiating with other members of the group. Also, a host may belong to more than one group at a time. Internet multicast can be implemented on top of a collection of networks that support hardware multicast (or broadcast) by extending the routing and forwarding functions implemented by the routers that connect these networks. This section de- scribes three such extensions: The first is based on distance-vector routing as described in Section 4.2.2; the second is based on link-state routing, described in Section 4.2.3; the third can build on any underlying routing protocol and is thus called Protocol Independent Multicast (PIM). Before looking at any of the multicast routing protocols, however, we need to look at the service model for IP multicast. We could imagine that a host wishing to send a packet to some number of internet hosts could enumerate all their addresses, but this would quickly become unscalable for large numbers of receivers—consider using the Internet to distribute a pay-per-view movie, for example. For this reason, IP multicast uses the idea of a multicast group that receivers may join. Each group has a specially assigned address, and senders to the group use that address as the destination address for their packets. In IPv4, these addresses are assigned in the class D address space, and IPv6 also has a portion of its address space reserved for multicast group addresses. Hosts join multicast groups using a protocol called Internet Group Management Protocol (IGMP). They use this to notify a router on their local network of their desire to receive packets sent to a certain multicast group. The protocols described below are concerned with how packets are distributed to the appropriate routers. Delivery of packets from the “last hop” router to the host is handled by the underlying multicast capability of the network, as described in Section 2.6. 332 4 Internetworking One perplexing question is how senders and receivers learn about multicast addresses. This is normally handled by out-of-band means, and there are some quite sophisticated tools to enable group addresses to be advertised on the Internet. 4.4.1 Link-State Multicast Adding multicast to a link-state routing algorithm is fairly straightforward, so we de- scribe it first. Recall that in link-state routing, each router monitors the state of its directly connected links and sends an update message to all of the other routers when- ever the state changes. Since each router receives enough information to reconstruct the entire topology of the network, it is able to use Dijkstra’s algorithm to compute the shortest-path spanning tree rooted at itself and reaching all possible destinations. The router uses this tree to determine the best next hop for each packet it forwards. All we have to do to extend this algorithm to support multicast is to add the set of groups that have members on a particular link (LAN) to the “state” for that link. The only question is how each router determines which groups have members on which links. As suggested in Section 3.2.3, the solution is to have each host periodically announce to the LAN the groups to which it belongs. The router simply monitors the LAN for such announcements. Should such announcements stop arriving after a period of time, the router then assumes that the host has left the group. Given full knowledge of which groups have members on which links, each router is able to compute the shortest-path multicast tree from any source to any group, again using Dijkstra’s algorithm. For example, given the internet illustrated in Figure 4.36, where the colored hosts belong to group G, the routers would compute the shortest-path multicast trees given in Figure 4.37 for sources A, B, and C. The routers would use these trees to decide how to forward packets addressed to multicast group G. For example, router R3 would forward a packet going from host A to group GtoR6. Keep in mind that each router must potentially keep a separate shortest-path multicast tree from every source to every group. This is obviously very expensive, so instead the router just computes and stores a cache of these trees—one for each source/group pair that is currently active. 4.4.2 Distance-Vector Multicast Adding multicast to the distance-vector algorithm is a bit trickier because the routers do not know the entire topology of the internet. Instead, recall that each router maintains a table of  Destination, Cost, NextHop  tuples, and exchanges a list of  Destination, Cost  pairs with its directly connected neighbors. Extending this algorithm to support multicast is a two-stage process. First, we need to design a broadcast mechanism that allows a packet to be forwarded to all the networks on the internet. Second, we need 4.4 Multicast 333 R1 R2 R3 R4 R5 R6 R7 A B C Figure 4.36 Example internet with members of group G in color. to refine this mechanism so that it prunes back networks that do not have hosts that belong to the multicast group. Reverse-Path Broadcast (RPB) Each router knows that the current shortest path to a given destination goes through NextHop. Thus, whenever it receives a multicast packet from source S, the router forwards the packet on all outgoing links (except the one on which the packet arrived) if and only if the packet arrived over the link that is on the shortest path to S (i.e., the packet came from the NextHop associated with S in the routing table). This strategy effectively floods packets outward from S, but does not loop packets back toward S. There are two major shortcomings to this approach. The first is that it truly floods the network; it has no provision for avoiding LANs that have no members in the multicast group. We address this problem in the next subsection. The second limitation is that a given packet will be forwarded over a LAN by each of the routers connected to that LAN. This is due to the forwarding strategy of flooding packets on all links other than the one on which the packet arrived, without regard to whether or not those links are part of the shortest-path tree rooted at the source. The solution to this second limitation is to eliminate the duplicate broadcast packets that are generated when more than one router is connected to a given LAN. 334 4 Internetworking R7R6 R5R4 R2 R2 B R6 R5R4 R1 R1 R1 A C R7 R6 R7 C Source Source Source R3 R5R4R3 R2 R3 C B A A B Figure 4.37 Example shortest-path multicast trees. 4.4 Multicast 335 One way to do this is to designate one router as the “parent” router for each link, relative to the source, where only the parent router is allowed to forward multicast packets from that source over the LAN. The router that has the shortest path to source S is selected as the parent; a tie between two routers would be broken according to which router has the smallest address. A given router can learn if it is the parent for the LAN (again relative to each possible source) based upon the distance-vector messages it exchanges with its neighbors. Notice that this refinement requires that each router keep, for each source, a bit for each of its incident links indicating whether or not it is the parent for that source/link pair. Keep in mind that in an internet setting, a “source” is a net- work, not a host, since an internet router is only interested in forwarding packets be- tween networks. The resulting mechanism is sometimes called reverse-path broadcast (RPB). Reverse-Path Multicast (RPM) RPB implements shortest-path broadcast. We now want to prune the set of networks that receives each packet addressed to group G to exclude those that have no hosts that are members of G. This can be accomplished in two stages. First, we need to recognize when a leaf network has no group members. Determining that a network is a leaf is easy—if the parent router as described in RPB is the only router on the network, then the network is a leaf. Determining if any group members reside on the network is accomplished by having each host that is a member of group G periodically announce this fact over the network, as described in our earlier description of link-state multicast. The router then uses this information to decide whether or not to forward a multicast packet addressed to G over this LAN. The second stage is to propagate this “no members of G here” information up the shortest-path tree. This is done by having the router augment the Destination, Cost pairs it sends to its neighbors with the set of groups for which the leaf network is interested in receiving multicast packets. This information can then be propagated from router to router, so that for each of its links, a given router knows for what groups it should forward multicast packets. Note that including all of this information in the routing update is a fairly expen- sive thing to do. In practice, therefore, this information is exchanged only when some source starts sending packets to that group. In other words, the strategy is to use RPB, which adds a small amount of overhead to the basic distance-vector algorithm, until a particular multicast address becomes active. At that time, routers that are not interested in receiving packets addressed to that group speak up, and that information is propagated to the other routers. 336 4 Internetworking 4.4.3 Protocol Independent Multicast (PIM) PIM was developed in response to the scaling problems of existing multicast routing protocols. In particular, it was recognized that the existing protocols did not scale well in environments where a relatively small proportion of routers want to receive traffic for a certain group. For example, broadcasting traffic to all routers until they explicitly ask to be removed from the distribution is not a good design choice if most routers don’t want to receive the traffic in the first place. This situation is sufficiently common that PIM divides the problem space into “sparse mode” and “dense mode.” Because the existing protocols were so poorly suited to the sparse environment, PIM sparse mode has received the most attention and is the focus of our discussion here. In PIM sparse mode (PIM-SM), routers explicitly join and leave the multicast group using PIM protocol messages known as Join and Prune messages. The question that arises is where to send those messages. To address this, PIM assigns a rendezvous point (RP) to each group. In general, a number of routers in a domain are configured to be candidate RPs, and PIM defines a set of procedures by which all the routers in a domain can agree on the router to use as the RP for a given group. These procedures are rather complex, as they must deal with a wide variety of scenarios, such as the failure of a candidate RP and the partitioning of a domain into two separate networks due to a number of link or node failures. For the rest of this discussion, we assume that all routers in a domain know the unicast IP address of the RP for a given group. A multicast forwarding tree is built as a result of routers sending Join messages to the RP. PIM-SM allows two types of trees to be constructed: a shared tree, which may be used by all senders, and a source-specific tree, which may be used only by a specific sending host. The normal mode of operation creates the shared tree first, followed by one or more source-specific trees if there is enough traffic to warrant it. Because building trees installs state in the routers along the tree, it is important that the default is to have only one tree for a group, not one for every sender to a group. When a router sends a Join message toward the RP for a group G, it is sent using normal IP unicast transmission. This is illustrated in Figure 4.38(a), in which router R4 is sending a Join to the rendezvous point for some group. The initial Join message is “wildcarded”; that is, it applies to all senders. A Join message clearly must pass through some sequence of routers before reaching the RP (e.g., R2). Each router along the path looks at the Join and creates a forwarding table entry for the shared tree, called a (*, G) entry (* meaning “all senders”). To create the forwarding table entry, it looks at the interface on which the Join arrived and marks that interface as one on which it should forward data packets for this group. It then determines which interface it will use to forward the Join toward the RP. This will be the only acceptable interface 4.4 Multicast 337 RP R3 R2 R4 R1 R5 (c) RP R3 R2 R4 R1 R5 (d) RP R3 R2 R4 R1 R5 (a) RP R3 R2 R4 R1 R5 (b) Join Join Join Join Join Rendezvous pointRP = Shared tree Source-specific tree for source R1 Figure 4.38 PIM operation: (a) R4 sends Join to RP and joins shared tree. (b) R5 joins shared tree. (c) RP builds source-specific tree to R1 by sending Join to R1. (d) R4 and R5 build source-specific tree to R1 by sending Joins to R1. for incoming packets sent to this group. It then forwards the Join toward the RP. Eventually, the message arrives at the RP, completing the construction of the tree branch. The shared tree thus constructed is shown as a colored line from the RP to R4 in Figure 4.38(a). As more routers send Joins toward the RP, they cause new branches to be added to the tree, as illustrated in Figure 4.38(b). Note that in this case, the Join only needs to travel to R2, which can add the new branch to the tree simply by adding a new outgoing interface to the forwarding table entry created for this group. R2 need not forward the Join on to the RP. Note also that the end result of this process is to build a tree whose root is the RP. 338 4 Internetworking RP R3 R2 R4 R1 R5 Host RP G RP G G G G G Figure 4.39 Delivery of a packet along a shared tree. R1 tunnels the packet to the RP, which forwards it along the shared tree to R4 and R5. At this point, suppose a host wishes to send a message to the group. To do so, it constructs a packet with the appropriate multicast group address as its destination and sends it to a router on its local network known as the designated router (DR). Suppose the DR is R1 in Figure 4.38. There is no state for this multicast group between R1 and the RP at this point, so instead of simply forwarding the multicast packet, R1 “tunnels” it to the RP. That is, R1 encapsulates the multicast packet inside a unicast IP packet that it sends to the unicast IP address of the RP. Just like a tunnel endpoint of the sort described in Section 4.1.8, the RP receives the packet addressed to it, looks at the payload of the unicast packet, and finds inside an IP packet addressed to the multicast address of this group. The RP, of course, does know what to do with such a packet—it sends it out onto the shared tree of which the RP is the root. In the example of Figure 4.38, this means that the RP sends the packet on to R2, which is able to forward it to R4 and R5. The complete delivery of a packet from R1 to R4 and R5 is shown in Figure 4.39. We see the tunneled packet travel from R1 to the RP with an extra IP header containing the unicast address of RP, and then the multicast packet addressed to G making its way along the shared tree to R4 and R5. At this point, we might be tempted to declare success, since all hosts can send to all receivers this way. However, there is some bandwidth inefficiency and processing cost in the encapsulation and decapsulation of packets on the way to the RP, so the RP has the option of forcing knowledge about this group into the intervening routers so that tunneling can be avoided. Its decision to exercise this option is based on the 4.4 Multicast 339 data rate of packets coming from a given source; only if this rate is high enough to warrant the effort will the RP take action. If it does, it sends a Join message toward the sending host (Figure 4.38(c)). As this Join travels toward the host, it causes the routers along the path (R3) to learn about the group, so that it will be possible for the DR to send the packet to the group as “native” (i.e., not tunneled) multicast packets. An important detail to note at this stage is that the Join message sent by the RP to the sending host is specific to that sender, whereas the previous ones sent by R4 and R5 applied to all senders. Thus the effect of the new Join is to create sender-specific state in the routers between the identified source and the RP. This is referred to as (S, G) state, since it applies to one sender to one group and contrasts with the (*, G) state that was installed between the receivers and the RP that applies to all senders. Thus, in Figure 4.38(c), we see a source-specific route from R1 to the RP (indicated by the dashed line) and a tree that is valid for all senders from the RP to the receivers (indicated by the colored line). The next possible optimization is to replace the entire shared tree with a source- specific tree. This is desirable because the path from sender to receiver via the RP might be significantly longer than the shortest possible path. This again is likely to be triggered by a high data rate being observed from some sender. In this case, the router at the downstream end of the tree—say, R4 in our example—sends a source-specific Join toward the source. As it follows the shortest path toward the source, the routers along the way create (S, G) state for this tree, and the result is a tree that has its root at the source, rather than the RP. Assuming both R4 and R5 made the switch to the source-specific tree, we would end up with the tree shown in Figure 4.38(d). Note that this tree no longer involves the RP at all. We have removed the shared tree from this picture to simplify the diagram, but in reality all routers with receivers for a group must stay on the shared tree in case new senders show up. We can now see why PIM is “protocol independent.” All of its mechanisms for building and maintaining trees depend on whatever unicast routing protocol is used in the domain. The formation of trees is entirely determined by the paths that Join messages follow, which is determined by the choice of shortest paths made by unicast routing. Thus, to be precise, PIM is “unicast routing protocol independent,” as compared to the other multicast routing protocols in this section, which are derived from either link-state or distance-vector routing. Note that PIM is very much bound up with the Internet Protocol—it is not protocol independent in terms of network-layer protocols. The design of PIM again illustrates the challenges in building scalable networks, and how scalability is sometimes pitted against some sort of optimality. The shared tree is certainly more scalable than a source-specific tree, in the sense that it reduces 340 4 Internetworking the total state in routers to be on the order of the number of groups rather than the number of senders times the number of groups. However, the source-specific tree is likely to be necessary to achieve efficient routing. 4.5 Multiprotocol Label Switching (MPLS) We conclude our discussion of IP by describing an idea that was originally viewed as a way to improve the performance of the Internet. The idea, called Multiprotocol Label Switching (MPLS), tries to combine some of the properties of virtual circuits with the flexibility and robustness of datagrams. On the one hand, MPLS is very much associated with the Internet Protocol’s datagram-based architecture—it relies on IP addresses and IP routing protocols to do its job. On the other hand, MPLS-enabled routers also forward packets by examining relatively short, fixed-length labels, and these labels have local scope, just like in a virtual circuit network. It is perhaps this marriage of two seemingly opposed technologies that has caused MPLS to have a somewhat mixed reception in the Internet engineering community. Before looking at how MPLS works, it is reasonable to ask, “What is it good for?” Many claims have been made for MPLS, but there are three main things that it is used for today: ■ To enable IP capabilities on devices that do not have the capability to forward IP datagrams in the normal manner ■ To forward IP packets along “explicit routes”—precalculated routes that don’t necessarily match those that normal IP routing protocols would select ■ To support certain types of virtual private network services It is worth noting that one of the original goals—improving performance— is not on the list. This has a lot to do with the advances that have been made in forwarding algorithms for IP routers in recent years, and with the complex set of factors beyond header processing that determine performance. The best way to understand how MPLS works is to look at some examples of its use. In the next three sections we will look at examples to illustrate the three applications of MPLS mentioned above. 4.5.1 Destination-Based Forwarding One of the earliest publications to introduce the idea of attaching labels to IP packets was a paper by Chandranmenon and Varghese that described an idea called “threaded indices.” A very similar idea is now implemented in MPLS-enabled routers. The fol- lowing example shows how this idea works. 4.5 Multiprotocol Label Switching (MPLS) 341 10.1.1/24 10.3.3/24 0 0 1 10.1.1 0 10.3.3 0 … Prefix Interface 10.1.1 10.3.3 0 … Prefix Interface 1 R2R1 R4 R3 Figure 4.40 Routing tables in example network. Consider the network in Figure 4.40. Each of the two routers on the far right (R3 and R4) has one connected network, with prefixes 10.1.1/24 and 10.3.3/24. The remaining routers (R1 and R2) have routing tables that indicate which outgoing inter- face each router would use when forwarding packets to one of those two networks. When MPLS is enabled on a router, the router allocates a label for each prefix in its routing table and advertises both the label and the prefix that it represents to its neighboring routers. This advertisement is carried in the “Label Distribution Protocol.” This is illustrated in Figure 4.41. Router R2 has allocated the label value 15 for the prefix 10.1.1 and the label value 16 for the prefix 10.3.3. These labels can be chosen at the convenience of the allocating router and can be thought of as indices into the routing table. After allocating the labels, R2 advertises the label bindings to its neighbors; in this case, we see R2 advertising a binding between the label 15 and the prefix 10.1.1 to R1. The meaning of such an advertisement is that R2 has said, in effect, “Please attach the label 15 to all packets sent to me that are destined to prefix 10.1.1.” R1 stores the label in a table alongside the prefix that it represents as the “remote” or “outgoing” label for any packets that it sends to that prefix. In Figure 4.41(c), we see another label advertisement from router R3 to R2 for the prefix 10.1.1, and R2 places the “remote” label that it learned from R3 in the appropriate place in its table. At this point, we can look at what happens when a packet is forwarded in this network. Suppose a packet destined to the IP address 10.1.1.5 arrives from the left to router R1. R1 in this case is referred to as a label edge router (LER); an LER performs a complete IP lookup on arriving IP packets, and then applies labels to them as a result of the lookup. In this case, R1 would see that 10.1.1.5 matches the prefix 10.1.1 in its 342 4 Internetworking 10.1.1/24 10.3.3/24 0 0 1 R2R1 R4 R3 10.1.1 0 10.3.3 0 … Prefix Interface 15 10.1.1 1 16 10.3.3 0 … Label Prefix Interface 15 10.1.1 1 16 10.3.3 0 … Label Prefix Interface Label = 15, Prefix = 10.1.1 10.1.1/24 10.3.3/2400 1 10.1.1 0 10.3.3 0 … Prefix Interface R2R1 R4 R3 Remote label 15 16 (a) (b) 10.1.1/24 10.3.3/24 0 0 1 R2R1 R4 R3 15 10.1.1 1 24 16 10.3.3 0 … Label Prefix Interface (c) 10. 1.1 0 10. 3.3 0 … Prefix Interface Remote label 15 16 Remote label Label = 24, Prefix = 10.1.1 Figure 4.41 (a) R2 allocates labels and advertises bindings to R1. (b) R1 stores the received labels in a table. (c) R3 advertises another binding, and R2 stores the received label in a table. 4.5 Multiprotocol Label Switching (MPLS) 343 forwarding table, and that this entry contains both an outgoing interface and a remote label value. R1 therefore attaches the remote label 15 to the packet before sending it. When the packet arrives at R2, R2 looks at the label in the packet. The forwarding table at R2 indicates that packets arriving with a label value of 15 should be sent out interface 1, and that they should carry the label value 24, as advertised by router R3. R2 therefore rewrites, or swaps, the label and forwards it to R3. What has been accomplished by all this application and swapping of labels? Observe that when R2 forwarded the packet in this example, it never actually needed to examine the IP address. Instead, R3 looked only at the incoming label. Thus, we have replaced the normal IP destination address lookup with a label lookup. To understand why this is significant, it helps to recall that although IP addresses are always the same length, IP prefixes are of variable length, and the IP destination address lookup algorithm needs to find the longest match—the longest prefix that matches the high- order bits in the IP address of the packet being forwarded. By contrast, the label forwarding mechanism just described is an exact match algorithm. It is possible to implement a very simple exact match algorithm, for example, by using the label as an index into an array, where each element in the array is one line in the forwarding table. Note that while the forwarding algorithm has been changed from longest match to exact match, the routing algorithm can be any standard IP routing algorithm (e.g., OSPF). The path that a packet will follow in this environment is the exact same path that it would have followed if MPLS were not involved—the path chosen by the IP routing algorithms. All that has changed is the forwarding algorithm. The major effect of changing the forwarding algorithm is that devices that nor- mally don’t know how to forward IP packets can be used in an MPLS network. The most notable early application of this result was to ATM switches, which can support MPLS without any changes to their forwarding hardware. ATM switches support the label swapping forwarding algorithm just described, and by providing these switches with IP routing protocols and a method to distribute label bindings, they could be turned into label switching routers (LSRs)—devices that run IP control protocols but use the label switching forwarding algorithm. More recently, the same idea has been applied to optical switches of the sort described in Section 3.1.2. Before we consider the purported benefits of turning an ATM switch into an LSR, we should tie up some loose ends. We have said that labels are “attached” to packets, but where exactly are they attached? The answer depends on the type of link on which packets are carried. Two common methods for carrying labels on packets are shown in Figure 4.42. When IP packets are carried as complete frames, as they are on most link types including Ethernet, token ring, and PPP, the label is inserted as a “shim” between the layer 2 header and the IP (or other layer 3) header, as shown in Figure 4.42(b). 344 4 Internetworking However, if an ATM switch is to function as an MPLS LSR, then the label needs to be in a place where the switch can use it, and that means it needs to be in the ATM cell header, exactly where we would nor- mally find the VCI and VPI fields, as shown in Figure 4.42(a). Having now devised a scheme by which an ATM switch can function as an LSR, what have we gained? One thing to note is that we could now build a net- work that used a mixture of conventional IP routers, label edge routers, and ATM switches functioning as LSRs, and they would all use the same routing protocols. To understand the benefits of using the same protocols, consider the alternative. In Figure 4.43(a) we see a set of routers intercon- nected by virtual circuits over an ATM net- work, a configuration called an “overlay” network. At one point in time, networks of this type were often built because commer- cially available ATM switches supported higher total throughput than routers. Today, networks like this are less common because routers have caught up with and even sur- passed ATM switches. However, these net- works still exist because of the significant installed base of ATM switches in network What Layer Is MPLS? There have been many debates about where MPLS belongs in the layered protocol architectures pre- sented in Section 1.3. Since the MPLS header is normally found between the layer 3 and the layer 2 headers in a packet, it is some- times referred to as a layer 2.5 pro- tocol. Some people argue that, since IP packets are encapsulated in- side MPLS headers, MPLS must be “below” IP, making it a layer 2 pro- tocol. Others argue that, since the control protocols for MPLS are, in large part, the same protocols as IP—MPLS uses IP routing proto- cols and IP addressing—then MPLS must be at the same layer as IP (i.e., layer 3). As we noted in Section 1.3, layered architectures are useful tools but they may not always exactly describe the real world, and MPLS is a good example of where strictly layerist views may be difficult to reconcile with reality. backbones, which in turn is partly a result of ATM’s ability to support a range of capabilities such as circuit emulation and virtual circuit services. In an overlay network, each router would potentially be connected to each of the other routers by a virtual circuit, but in this case for clarity we have just shown the circuits from R1 to all of its peer routers. R1 has five routing neighbors and needs to exchange routing protocol messages with all of them—we say that R1 has five routing adjacencies. By contrast, in Figure 4.43(b), the ATM switches have been replaced with LSRs. There are no longer virtual circuits interconnecting the routers. Thus R1 has only one adjacency, with LSR1. In large networks, running MPLS on the switches leads to a significant reduction in the number of adjacencies that each router must maintain 4.5 Multiprotocol Label Switching (MPLS) 345 Label headerPPP header Layer 3 header “Shim” header (for PPP, Ethernet, etc.) ATM cell header HEC Label DATACLPPTIVCIGFC VPI(a) (b) Figure 4.42 (a) Label on an ATM-encapsulated packet. (b) Label on a frame- encapsulated packet. R3 R4 R6 R5 R1 R2 LSR1 LSR2 LSR3 R3 R4 R6 R2 R1 R5 (a) (b) Figure 4.43 (a) Routers connect to each other using an “overlay”of virtual circuits. (b) Routers peer directly with LSRs. 346 4 Internetworking and can greatly reduce the amount of work that the routers have to do to keep each other informed of topology changes. A second benefit of running the same routing protocols on edge routers and on the LSRs is that the edge routers now have a full view of the topology of the network. This means that if some link or node fails inside the network, the edge routers will have a better chance of picking a good new path than if the ATM switches rerouted the affected VCs without the knowledge of the edge routers. Note that the step of “replacing” ATM switches with LSRs is actually achieved by changing the protocols running on the switches, but typically no change to the forwarding hardware is needed. That is, an ATM switch can often be converted to an MPLS LSR by upgrading only its software. Furthermore, an MPLS LSR might continue to support standard ATM capabilities at the same time as it runs the MPLS control protocols. More recently, the idea of running IP control protocols on devices that are unable to forward IP packets natively has been extended to optical switches and TDM devices such as SONET multiplexers. This is known as generalized MPLS (GMPLS). Part of the motivation for GMPLS was to provide routers with topological knowledge of an optical network, just as in the ATM case. Even more important was the fact that there were no standard protocols for controlling optical devices, and so MPLS seemed like a natural fit for that job. 4.5.2 Explicit Routing In Section 3.1.3 we introduced the concept of source routing. IP has a source routing option, but it is not widely used for several reasons, including the fact that only a limited number of hops can be specified, and because it is usually processed outside the “fast path” on most routers. MPLS provides a convenient way to add capabilities similar to source routing to IP networks, although the capability is more often called “explicit routing” rather than “source routing.” One reason for the distinction is that it usually isn’t the real source of the packet that picks the route. More often it is one of the routers inside a service provider’s network. Figure 4.44 shows an example of how the explicit routing capability of MPLS might be applied. This sort of network is often called a “fish” network because of its shape (the routers R1 and R2 form the tail; R7 is at the head). Suppose that the operator of the network in Figure 4.44 has determined that any traffic flowing from R1 to R7 should follow the path R1-R3-R6-R7, and that any traffic going from R2 to R7 should follow the path R2-R3-R4-R7. One reason for such a choice would be to make good use of the capacity available along the two distinct paths from R3 to R7. This cannot easily be accomplished with normal IP routing because R3 doesn’t look at where traffic came from in making its forwarding decisions. 4.5 Multiprotocol Label Switching (MPLS) 347 R1 R3 R2 R6 R4 R5 R7 Figure 4.44 A network requiring explicit routing. Because MPLS uses label swapping to forward packets, it is easy enough to achieve the desired routing if the routers are MPLS-enabled. If R1 and R2 attach distinct labels to packets before sending them to R3, then R3 can forward packets from R1 and R2 along different paths. The question that then arises is, How do all the routers in the network agree on what labels to use and how to forward packets with particular labels? Clearly, we can’t use the same procedures as described in the preceding section to distribute labels because those procedures establish labels that cause packets to follow the normal paths picked by IP routing, which is exactly what we are trying to avoid. Instead, a new mechanism is needed. It turns out that the protocol used for this task is the Resource Reservation Protocol (RSVP). We’ll talk more about this protocol in Section 6.5.2, but for now it suffices to say that it is possible to send an RSVP message along an explicitly specified path (e.g., R1-R3-R6-R7) and use it to set up label forwarding table entries all along that path. This is very similar to the process of establishing a virtual circuit described in Section 3.1.2. One of the applications of explicit routing is “traffic engineering,” which refers to the task of ensuring that sufficient resources are available in a network to meet the demands placed on it. Controlling exactly which paths the traffic flows on is an important part of traffic engineering. Explicit routing can also help to make networks more resilient in the face of failure, using a capability called fast reroute. For example, it is possible to precalculate a path from router A to router B that explicitly avoids a certain link L. In the event that link L fails, router A could send all traffic destined to B down the precalculated path. The combination of precalculation of the “backup path” and the explicit routing of packets along the path means that A doesn’t need to wait for routing protocol packets to make their way across the network or for routing algorithms to be executed by various other nodes in the network. In certain circumstances, this can significantly reduce the time taken to reroute packets around a point of failure. 348 4 Internetworking One final point to note about explicit routing is that explicit routes need not be calculated by a network operator as in the above example. There are a range of algorithms that routers can use to calculate explicit routes automatically. The most common of these is called constrained shortest path first (CSPF), which is like the link-state algorithms described in Section 4.2.3, but which also takes “constraints” into account. For example, if it was required to find a path from R1 to R7 that could carry an offered load of 100 Mbps, we could say that the “constraint” is that each link must have at least 100 Mbps of available capacity. CSPF addresses this sort of problem. More details on CSPF, and the applications of explicit routing, are provided in the “Further Reading” section at the end of the chapter. 4.5.3 Virtual Private Networks and Tunnels We first talked about virtual private networks (VPNs) in Section 4.1.8, and we noted that one way to build them was using tunnels. It turns out that MPLS can be thought of as a way to build tunnels, and this makes it suitable for building VPNs of various types. The simplest form of MPLS VPN to understand is a “layer 2” VPN. In this type of VPN, MPLS is used to tunnel layer 2 data (such as Ethernet frames or ATM cells) across a network of MPLS-enabled routers. Recall from Section 4.1.8 that one reason for tunnels is to provide some sort of network service (such as multicast) that is not supported by some routers in the network. The same logic applies here: IP routers are not ATM switches, so you cannot provide an ATM virtual circuit service across a network of conventional routers. However, if you had a pair of routers interconnected by a tunnel, they could send ATM cells across the tunnel and emulate an ATM circuit. The term for this technique within the IETF is pseudowire emulation. Figure 4.45 illustrates the idea. We have already seen how IP tunnels are built: The router at the entrance of the tunnel wraps the data to be tunneled in an IP header (the “tunnel header”), which R2 Tail R3 Head ATM cells arrive Cells sent into tunnel at head Tunneled data arrives at tail ATM cells sent Figure 4.45 An ATM circuit is emulated by a tunnel. 4.5 Multiprotocol Label Switching (MPLS) 349 represents the address of the router at the far end of the tunnel, and sends the data like any other IP packet. The receiving router receives the packet with its own address in the header, strips the tunnel header, and finds the data that was tunneled, which it then processes. Exactly what it does with that data depends on what it is. For example, if it were another IP packet, it would then be forwarded like a normal IP packet. However, it need not be an IP packet, as long as the receiving router knows what to do with non-IP packets. We’ll return to the issue of how to handle non-IP data in a moment. An MPLS tunnel is not too different from an IP tunnel, except that the “tunnel header” consists of an MPLS header rather than an IP header. Looking back to our first example, in Figure 4.41, we saw that router R1 attached a label (15) to every packet that it sent toward prefix 10.1.1. Such a packet would then follow the path R1-R2-R3, with each router in the path examining only the MPLS label. Thus, we observe that there was no requirement that R1 only send IP packets along this path— any data could be wrapped up in the MPLS header, and it would follow the same path because the intervening routers never look beyond the MPLS header. In this regard, an MPLS header is just like an IP tunnel header.7 The only issue with sending non-IP traffic along a tunnel, MPLS or otherwise, is this: What do we do with non-IP traffic when it reaches the end of the tunnel? The general solution is to carry some sort of demultiplexing identifier in the tunnel payload that tells the router at the end of the tunnel what to do. It turns out that an MPLS label is a perfect fit for such an identifier. An example will make this clear. Let’s assume we want to tunnel ATM cells from one router to another across a network of MPLS-enabled routers, as in Figure 4.45. Further, we assume that the goal is to emulate an ATM virtual circuit; that is, cells arrive at the entrance, or head, of the tunnel on a certain input port with a certain VCI and should leave the tail end of the tunnel on a certain output port and potentially different VCI. This can be accomplished by configuring the “head” and “tail” routers as follows: ■ The head router needs to be configured with the incoming port, the incoming VCI, the “demultiplexing label” for this emulated circuit, and the address of the tunnel end router. ■ The tail end router needs to be configured with the outgoing port, the outgoing VCI, and the demultiplexing label. Once the routers are provided with this information, we can see how an ATM cell would be forwarded. Figure 4.46 illustrates the steps. 7Note, however, that an MPLS header is only 4 bytes long, compared to 20 for an IP header, which implies a bandwidth savings when MPLS is used. 350 4 Internetworking R2 Tail R3 Head 1. ATM cells arrive 6. ATM cells sent101 202 101DL 101DLTL 101DL 101DLTL 2. Demux label added 4. Packet is forwarded to tail3. Tunnel label added 5. Demux label examined Figure 4.46 Forwarding ATM cells along a tunnel. 1 An ATM cell arrives on the designated input port with the appropriate VCI value (101 in this example). 2 The head router attaches the demultiplexing label that identifies the emulated circuit. 3 The head router then attaches a second label, which is the tunnel label that will get the packet to the tail router. This label is learned by mechanisms just like those described in Section 4.5.1. 4 Routers between the head and tail forward the packet using only the tunnel label. 5 The tail router removes the tunnel label, finds the demultiplexing label, and recognizes the emulated circuit. 6 The tail router modifies the ATM VCI to the correct value (202 in this case) and sends it out the correct port. One item in this example that might be surprising is that the packet has two labels attached to it. This is one of the interesting features of MPLS—labels may be “stacked” on a packet to any depth. This provides some useful scaling capabilities. In this example, it enables a single tunnel to carry a potentially large number of emulated circuits. The same techniques described here can be applied to emulate many other layer 2 services, including Frame Relay and Ethernet. It is worth noting that virtually identical capabilities can be provided using IP tunnels; the main advantage of MPLS here is the shorter tunnel header. Before MPLS was used to tunnel layer 2 services, it was also being used to support layer 3 VPNs. We won’t go into the details of layer 3 VPNs, which are quite complex— 4.5 Multiprotocol Label Switching (MPLS) 351 Provider network VPN A/Site 3VPN A/Site 1 VPN A/Site 2 VPN B/Site 2 VPN B/Site 1 VPN B/Site 3 Figure 4.47 Example of a layer 3 VPN. Customers A and B each obtain a virtually private IP service from a single provider. see the “Further Reading” section for some good sources of more information—but we will note that they represent one of the most popular uses of MPLS today. Layer 3 VPNs also use stacks of MPLS labels to tunnel packets across an IP network. However, the packets that are tunneled are themselves IP packets—hence the name “layer 3 VPNs.” In a layer 3 VPN, a single service provider operates a network of MPLS- enabled routers and provides a “virtually private” IP network service to any number of distinct customers. That is, each customer of the provider has some number of sites, and the service provider creates the illusion for each customer that there are no other customers on the network. The customer sees an IP network interconnecting his own sites, and no other sites. This means that each customer is isolated from all other customers in terms of both routing and addressing. Customer A can’t send packets directly to customer B, and vice versa.8 Customer A can even use IP addresses that have also been used by customer B. The basic idea is illustrated in Figure 4.47. As in 8Customer A in fact usually can send data to customer B in some restricted way. Most likely, both customer A and customer B have some connection to the global Internet, and thus it is probably possible for customer A to send email messages, for example, to the mail server inside customer B’s network. The “privacy” offered by a VPN prevents customer A from having unrestricted access to all the machines and subnets inside customer B’s network. 352 4 Internetworking layer 2 VPNs, MPLS is used to tunnel packets from one site to another. However, the configuration of the tunnels is performed automatically by some fairly elaborate use of BGP, which is beyond the scope of this book. In summary, MPLS is a rather versatile tool that has been applied to a wide range of different networking problems. It combines the label swapping forwarding mechanism that is normally associated with virtual circuit networks with the routing and control protocols of IP datagram networks to produce a class of network that is somewhere between the two conventional extremes. This extends the capabilities of IP networks to enable, among other things, more precise control of routing and the support of a range of VPN services. 4.6 Summary The main theme of this chapter was how to build big networks by interconnecting smaller networks. We looked at bridging in the last chapter, but it is a technique that is mostly used to interconnect a small to moderate number of similar networks. What bridging does not do well is tackle the two closely related problems of building very large networks: heterogeneity and scale. The Internet Protocol is the key tool for dealing with these problems, and it provided most of the examples for this chapter. IP tackles heterogeneity by defining a simple, common service model for an internetwork, which is based on the best-effort delivery of IP datagrams. An important part of the service model is the global addressing scheme, which enables any two nodes in an internetwork to uniquely identify each other for the purposes of exchanging data. The IP service model is simple enough to be supported by any known networking technology, and the ARP mechanism is used to translate global IP addresses into local link-layer addresses. A crucial aspect of the operation of an internetwork is the determination of efficient routes to any destination in the internet. Internet routing algorithms solve this problem in a distributed fashion; this chapter introduced the two major classes of algorithms—link-state and distance-vector—along with examples of their application (RIP and OSPF). We also examined the extensions to IP routing that will support mobile hosts. We then saw a succession of scaling problems and the ways that IP deals with them. The major scaling issues are the efficient use of address space and the growth of routing tables as the Internet grows. The hierarchical IP address format, with its network and host parts, gives us one level of hierarchy to manage scale. Subnetting lets us make more efficient use of network numbers and helps consolidate routing information; in effect, it adds one more level of hierarchy to the address. Classless routing (CIDR) lets us introduce more levels of hierarchy and achieve further routing Open Issue: Deployment of IPv6 353 aggregation. Autonomous systems allow us to partition the routing problem into two parts, interdomain and intradomain routing, each of which is much smaller than the total routing problem would be. These mechanisms have enabled today’s Internet to sustain remarkable growth. Eventually, all of these mechanisms will be unable to keep up with the Internet’s growth, and a new address format will be needed. This will require a new IP datagram format and a new version of the protocol. Originally known as IP Next Generation (IPng), this new protocol is now known as IPv6, and it provides a 128-bit address with CIDR-like addressing and routing. While many new capabilities have been claimed for IPv6, its main advantage remains its ability to support an extremely large number of addressable devices. More than 10 years have elapsed since the shortage of IPv4 address space became serious enough to war- rant proposals for a new version of IP. The orginal IPv6 specification is now more than 7 years old. IPv6- capable host operating systems are OPEN ISSUE Deployment of IPv6 now widely available and the major router vendors offer varying degrees of support for IPv6 in their products. Yet the deployment of IPv6 in the Internet has not, at the time of writing, begun in any meaningful way. It is worth wondering when deployment is likely to begin in earnest, and what will cause it. One reason why IPv6 has not been needed sooner is because of the extensive use of NAT (network address translation, described earlier in this chapter). As providers viewed IPv4 addresses as a scarce resource, they handed out fewer of them to their customers or charged for the number of addresses used; customers responded by hiding many of their devices behind a NAT box and a single IPv4 address. For example, it is likely that most home networks with more than one IP-capable device have some sort of NAT in the network to conserve addresses. So one factor that might drive IPv6 deployment would be applications that don’t work well with NAT. While client/server applications work reasonably well when the client’s address is “hidden” behind a NAT box, peer-to-peer applications fare less well. Examples of applications that would work better without NAT and would therefore benefit from more liberal address allocation policies are multiplayer gaming and IP telephony. Obtaining blocks of IPv4 addresses has been getting more difficult for years, and this is particularly noticeable in countries outside the United States. As the difficulty increases, the incentive for providers to start offering IPv6 addresses to their customers 354 4 Internetworking also rises. At the same time, for existing providers, offering IPv6 is a substantial addi- tional cost because they don’t get to stop supporting IPv4 when they start to offer IPv6. This means, for example, that the size of a provider’s routing tables can only increase initially because they need to carry all the existing IPv4 prefixes plus new IPv6 prefixes. At the moment, IPv6 deployment is happening almost exclusively in research networks. A few service providers are starting to offer it (often with some incentive from national governments). It seems hard to imagine that the Internet can continue to grow indefinitely without IPv6 seeing some more significant deployments, but it also seems likely that the overwhelming majority of hosts and networks will be IPv4-only for the foreseeable future. FURTHER READING Not surprisingly, there have been countless papers written on various aspects of the Internet. Of these, we recommend two as must reading: The paper by Cerf and Kahn is the one that originally introduced the TCP/IP architecture and is worth reading just for its historical perspective; the paper by Bradner and Mankin gives an informative overview of how the rapidly growing Internet has stressed the scalability of the orig- inal architecture, ultimately resulting in the next-generation IP. The paper by Paxson describes a study of how routers behave in the Internet. It also happens to be a good example of how researchers are now studying the dynamic behavior of the Internet. The final paper discusses multicast, presenting the approach to multicast originally used on the MBone. ■ Cerf, V., and R. Kahn. A protocol for packet network intercommunica- tion. IEEE Transactions on Communications COM-22(5):637–648, May 1974. ■ Bradner, S., and A. Mankin. The recommendation for the next generation IP protocol. Request for Comments 1752, January 1995. ■ Paxson, V. End-to-end routing behavior in the Internet. SIGCOMM ’96, pages 25–38, August 1996. ■ Deering, S., and D. Cheriton. Multicast routing in datagram internetworks and extended LANs. ACM Transactions on Computer Systems 8(2):85–110, May 1990. Beyond these papers, Perlman gives an excellent explanation of routing in an internet, including coverage of both bridges and routers [Per00]. Also, the book by Lynch and Rose gives general information on the scalability of the Internet [Cha93]. Exercises 355 Some interesting experimental studies of the behavior of Internet routing are presented in Labovitz et al. [LAAJ00]. Many of the techniques and protocols developed to help the Internet scale are described in RFCs: Subnetting is described in Mogul and Postel [MP85], CIDR is described in Fuller et al. [FLYV93], RIP is defined in Hedrick [Hed88] and Malkin [Mal93], OSPF is defined in Moy [Moy98], and BGP-4 is defined in Rekhter and Li [RL95]. The OSPF specification, at over 200 pages, is one of the longer RFCs around, but also contains an unusual wealth of detail about how to implement a protocol. A collection of RFCs related to IPv6 can be found in Bradner and Mankin [BM95], and the most recent IPv6 spec is by Deering and Hinden [DH98]. The reasons to avoid IP fragmentation are examined in Kent and Mogul [KM87], and the path MTU discovery technique is described in Mogul and Deering [MD90]. Protocol Independent Multicast (PIM) is described in Deering et al. [DEF+96] and [EFH+98]. There has been a lot of work developing algorithms that can be used by routers to do fast lookup of IP addresses. (Recall that the problem is that the router needs to match the longest prefix in the forwarding table.) PATRICIA trees are one of the first algo- rithms applied to this problem [Mor68]. More recent work is reported in [DBCP97], [WVTP97], [LS98], and [SVSW98]. For an overview of how these algorithms can be used to build a high-speed router, see Partridge et al. [Par98]. MPLS and the related protocols that fed its development are described in Chan- dranmenon and Varghese [CV95], Rekhter et al. [RDR+97], and Davie and Rekhter [DR00]. The latter reference describes many applications of MPLS such as traffic engineering, fast recovery from network failures, and virtual private networks. [RR99] provides the specification of MPLS/BGP VPNs, a form of layer 3 VPN that can be pro- vided over MPLS networks. Finally, we recommend the following live references: ■ http://www.ietf.org: the IETF home page, from which you can get RFCs, In- ternet drafts, and working group charters ■ http://playground.sun.com/pub/ipng/html/ipng-main.html:current state of IPv6 EXERCISES 1 What aspect of IP addresses makes it necessary to have one address per network interface, rather than just one per host? In light of your answer, why does IP tolerate point-to-point interfaces that have nonunique addresses or no addresses? 2 Why does the Offset field in the IP header measure the offset in 8-byte units? (Hint: Recall that the Offset field is 13 bits long.) 356 4 Internetworking 3 Some signalling errors can cause entire ranges of bits in a packet to be overwritten by all 0s or all 1s. Suppose all the bits in the packet including the Internet checksum are overwritten. Could a packet with all 0s or all 1s be a legal IPv4 packet? Will the Internet checksum catch that error? Why or why not? 4 Suppose a TCP message that contains 2048 bytes of data and 20 bytes of TCP header is passed to IP for delivery across two networks of the Internet (i.e., from the source host to a router to the destination host). The first network uses 14-byte headers and has an MTU of 1024 bytes; the second uses 8-byte headers with an MTU of 512 bytes. Each network’s MTU gives the size of the largest IP datagram that can be carried in a link-layer frame. Give the sizes and offsets of the sequence of fragments delivered to the network layer at the destination host. Assume all IP headers are 20 bytes. 5 Path MTU is the smallest MTU of any link on the current path (route) between two hosts. Assume we could discover the path MTU of the path used in the previous exercise, and that we use this value as the MTU for all the path segments. Give the sizes and offsets of the sequence of fragments delivered to the network layer at the destination host. 6 Suppose an IP packet is fragmented into 10 fragments, each with a 1% (inde- pendent) probability of loss. To a reasonable approximation, this means there is a 10% chance of losing the whole packet due to loss of a fragment. What is the probability of net loss of the whole packet if the packet is transmitted twice, (a) assuming all fragments received must have been part of the same transmission? (b) assuming any given fragment may have been part of either transmission? (c) Explain how use of the Ident field might be applicable here. 7 Suppose the fragments of Figure 4.5(b) all pass through another router onto a link with an MTU of 380 bytes, not counting the link header. Show the fragments produced. If the packet were originally fragmented for this MTU, how many fragments would be produced? 8 What is the maximum bandwidth at which an IP host can send 576-byte packets without having the Ident field wrap around within 60 seconds? Suppose IP’s max- imum segment lifetime (MSL) is 60 seconds; that is, delayed packets can arrive up to 60 seconds late but no later. What might happen if this bandwidth were exceeded? Exercises 357 9 ATM AAL3/4 uses fields Btag/Etag, BASize/Len, Type, SEQ, MID, Length, and CRC-10 to implement fragmentation into cells. IPv4 uses Ident, Offset, and the M bit in Flags, among others. What is the IP analog, if any, for each AAL3/4 field? Does each IP field listed here have an AAL3/4 analog? How well do these fields correspond? 10 Why do you think IPv4 has fragment reassembly done at the endpoint, rather than at the next router? Why do you think IPv6 abandoned fragmentation entirely? Hint: Think about the differences between IP-layer fragmentation and link-layer fragmentation. 11 Having ARP table entries time out after 10–15 minutes is an attempt at a reason- able compromise. Describe the problems that can occur if the timeout value is too small or too large. 12 IP currently uses 32-bit addresses. If we could redesign IP to use the 6-byte MAC address instead of the 32-bit address, would we be able to eliminate the need for ARP? Explain why or why not. 13 Suppose hosts A and B have been assigned the same IP address on the same Eth- ernet, on which ARP is used. B starts up after A. What will happen to A’s existing connections? Explain how “self-ARP” (querying the network on startup for one’s own IP address) might help with this problem. 14 Suppose an IP implementation adheres literally to the following algorithm on receipt of a packet, P, destined for IP address D: if (Ethernet address for D is in ARP cache) send Pelse send out an ARP query for D put P into a queue until the response comes back (a) If the IP layer receives a burst of packets destined for D, how might this algorithm waste resources unnecessarily? (b) Sketch an improved version. (c) Suppose we simply drop P, after sending out a query, when cache lookup fails. How would this behave? (Some early ARP implementations allegedly did this.) 358 4 Internetworking 2 3 6 2 8 1 D A F E B C Figure 4.48 Network for Exercises 15, 17, and 20. AB E DCF 21 23 523 Figure 4.49 Network for Exercise 16. 15 For the network given in Figure 4.48, give global distance-vector tables like those of Tables 4.5 and 4.8 when (a) each node knows only the distances to its immediate neighbors. (b) each node has reported the information it had in the preceding step to its immediate neighbors. (c) step (b) happens a second time. 16 For the network given in Figure 4.49, give global distance-vector tables like those of Tables 4.5 and 4.8 when (a) each node knows only the distances to its immediate neighbors. (b) each node has reported the information it had in the preceding step to its immediate neighbors. (c) step (b) happens a second time. 17 For the network given in Figure 4.48, show how the link-state algorithm builds the routing table for node D. Exercises 359 AF Node Cost NextHop Node Cost NextHop B 1 BA3 E C 2 BB2 C D 1 DC1 C E 2 BD2 E F 3 DE1 E Table 4.12 Forwarding tables for Exercise 18. AF Node Cost NextHop Node Cost NextHop B 1 BA2 C C 1 CB3 C D 2 BC1 C E 3 CD2 C F 2 CE1 E Table 4.13 Forwarding tables for Exercise 19. 18 Suppose we have the forwarding tables shown in Table 4.12 for nodes A and F, in a network where all links have cost 1. Give a diagram of the smallest network consistent with these tables. 19 Suppose we have the forwarding tables shown in Table 4.13 for nodes A and F, in a network where all links have cost 1. Give a diagram of the smallest network consistent with these tables. 20 For the network in Figure 4.48, suppose the forwarding tables are all established as in Exercise 15 and then the C–E link fails. Give (a) the tables of A, B, D, and F after C and E have reported the news. (b) the tables of A and D after their next mutual exchange. (c) the table of C after A exchanges with it. 360 4 Internetworking SubnetNumber SubnetMask NextHop 128.96.39.0 255.255.255.128 Interface 0 128.96.39.128 255.255.255.128 Interface 1 128.96.40.0 255.255.255.128 R2 192.4.153.0 255.255.255.192 R3 Default R4 Table 4.14 Routing table for Exercise 21. SubnetNumber SubnetMask NextHop 128.96.170.0 255.255.254.0 Interface 0 128.96.168.0 255.255.254.0 Interface 1 128.96.166.0 255.255.254.0 R2 128.96.164.0 255.255.252.0 R3 Default R4 Table 4.15 Routing table for Exercise 22. 21 Suppose a router has built up the routing table shown in Table 4.14. The router can deliver packets directly over interfaces 0 and 1, or it can forward packets to routers R2, R3, or R4. Describe what the router does with a packet addressed to each of the following destinations: (a) 128.96.39.10 (b) 128.96.40.12 (c) 128.96.40.151 (d) 192.4.153.17 (e) 192.4.153.90 22 Suppose a router has built up the routing table shown in Table 4.15. The router can deliver packets directly over interfaces 0 and 1, or it can forward packets to routers R2, R3, or R4. Assume the router does the longest prefix match. Describe Exercises 361 EAB Figure 4.50 Simple network for Exercise 23. what the router does with a packet addressed to each of the following destinations: (a) 128.96.171.92 (b) 128.96.167.151 (c) 128.96.163.151 (d) 128.96.169.192 (e) 128.96.165.121 23 Consider the simple network in Figure 4.50, in which A and B exchange distance- vector routing information. All links have cost 1. Suppose the A–E link fails. (a) Give a sequence of routing table updates that leads to a routing loop between A and B. (b) Estimate the probability of the scenario in (a), assuming A and B send out routing updates at random times, each at the same average rate. (c) Estimate the probability of a loop forming if A broadcasts an updated report within 1 second of discovering the A–E failure, and B broadcasts every 60 seconds uniformly. 24 Consider the situation involving the creation of a routing loop in the network of Figure 4.15 when the A–E link goes down. List all sequences of table updates among A, B, and C, pertaining to destination E, that lead to the loop. Assume that table updates are done one at a time, that the split horizon technique is observed by all participants, and that A sends its initial report of E’s unreachability to B before C. You may ignore updates that don’t result in changes. 25 Suppose a set of routers all use the split horizon technique; we consider here under what circumstances it makes a difference if they use poison reverse in addition. (a) Show that poison reverse makes no difference in the evolution of the routing loop in the two examples described in Section 4.2.2, given that the hosts involved use split horizon. (b) Suppose split horizon routers A and B somehow reach a state in which they forward traffic for a given destination X toward each other. Describe how this situation will evolve with and without the use of poison reverse. 362 4 Internetworking EAB D E ABand Figure 4.51 Networks for Exercise 26. A C BFG Figure 4.52 Network for Exercise 27. (c) Give a sequence of events that leads A and B to a looped state as in (b), even if poison reverse is used. Hint: Suppose B and A connect through a very slow link. They each reach X through a third node, C, and simultaneously advertise their routes to each other. 26 Hold-down is another distance-vector loop-avoidance technique, whereby hosts ignore updates for a period of time until link failure news has had a chance to propagate. Consider the networks in Figure 4.51, where all links have cost 1, except E–D with cost 10. Suppose that the E–A link breaks and B reports its loop- forming E route to A immediately afterward (this is the false route, via A). Specify the details of a hold-down interpretation, and use this to describe the evolution of the routing loop in both networks. To what extent can hold down prevent the loop in the EAB network without delaying the discovery of the alternative route in the EABD network? 27 Consider the network in Figure 4.52, using link-state routing. Suppose the B–F link fails, and the following then occur in sequence: (a) Node H is added to the right side with a connection to G. (b) Node D is added to the left side with a connection to C. (c) A new link D–A is added. The failed B–F link is now restored. Describe what link-state packets will flood back and forth. Assume that the initial sequence number at all nodes is 1, and Exercises 363 ABC DE5 5 4 1222 Figure 4.53 Network for Exercise 28. ADE BC6 5 3 1131 Figure 4.54 Network for Exercise 29. A B C Figure 4.55 Network for Exercise 30. that no packets time out, and that both ends of a link use the same sequence number in their LSP for that link, greater than any sequence number either used before. 28 Give the steps as in Table 4.9 in the forward search algorithm as it builds the routing database for node A in the network shown in Figure 4.53. 29 Give the steps as in Table 4.9 in the forward search algorithm as it builds the routing database for node A in the network shown in Figure 4.54. 30 Suppose that nodes in the network shown in Figure 4.55 participate in link-state routing, and C receives contradictory LSPs: One from A arrives claiming the A–B link is down, but one from B arrives claiming the A–B link is up. (a) How could this happen? (b) What should C do? What can C expect? Do not assume that LSPs contain any synchronized timestamp. 364 4 Internetworking A 1 2 3 B 4 Provider Q Provider P Provider R Figure 4.56 Network for Exercise 31. 31 Consider the network shown in Figure 4.56, in which horizontal lines represent transit providers and numbered vertical lines are interprovider links. (a) How many routes to P could provider Q’s BGP speakers receive? (b) Suppose Q and P adopt the policy that outbound traffic is routed to the closest link to the destination’s provider, thus minimizing their own cost. What paths will traffic from host A to host B and from host B to host A take? (c) What could Q do to have the B−→ A traffic use the closer link 1? (d) What could Q do to have the B−→ A traffic pass through R? 32 Give an example of an arrangement of routers grouped into autonomous systems so that the path with the fewest hops from a point A to another point B crosses the same AS twice. Explain what BGP would do with this situation. 33 Let A be the number of autonomous systems on the Internet, and let D (for diameter) be the maximum AS path length. (a) Give a connectivity model for which Dis of order log Aand another for which D is of order √A. (b) Assuming each AS number is 2 bytes and each network number is 4 bytes, give an estimate for the amount of data a BGP speaker must receive to keep track of the AS path to every network. Express your answer in terms of A, D, and the number of networks N. 34 Suppose IP routers learned about IP networks and subnets the way Ethernet learn- ing bridges learn about hosts: by noting the appearance of new ones and the in- terface by which they arrive. Compare this with existing distance-vector router learning (a) for a leaf site with a single attachment to the Internet, and (b) for internal use at an organization that did not connect to the Internet. Exercises 365 Rest of Internet A CR1RBR2 D B Figure 4.57 Site for Exercise 39. Assume that routers only receive new-network notices from other routers, and that the originating routers receive their IP network information via configuration. 35 IP hosts that are not designated routers are required to drop packets misaddressed to them, even if they would otherwise be able to forward them correctly. In the absence of this requirement, what would happen if a packet addressed to IP address A were inadvertently broadcast at the link layer? What other justifications for this requirement can you think of? 36 Read the man page or other documentation for the Unix/Windows utility netstat. Use netstat to display the current IP routing table on your host. Explain the purpose of each entry. What is the practical minimum number of entries? 37 Use the Unix utility traceroute (Windows tracert) to determine how many hops it is from your host to other hosts in the Internet (e.g., cs.princeton.edu or www.cisco.com). How many routers do you traverse just to get out of your local site? Read the man page or other documentation for traceroute and explain how it is implemented. 38 What will happen if traceroute is used to find the path to an unassigned address? Does it matter if the network portion or only the host portion is unassigned? 39 A site is shown in Figure 4.57. R1 and R2 are routers; R2 connects to the outside world. Individual LANs are Ethernets. RB is a bridge router; it routes traffic addressed to it and acts as a bridge for other traffic. Subnetting is used inside the site; ARP is used on each subnet. Unfortunately, host A has been misconfigured and doesn’t use subnets. Which of B, C, D can A reach? 366 4 Internetworking C BA Figure 4.58 Network for Exercise 41. 40 An organization has a class C network 200.1.1 and wants to form subnets for four departments, with hosts as follows: A 72 hosts B 35 hosts C 20 hosts D 18 hosts There are 145 hosts in all. (a) Give a possible arrangement of subnet masks to make this possible. (b) Suggest what the organization might do if department D grows to 34 hosts. 41 Suppose hosts A and B are on an Ethernet LAN with class C IP network address 200.0.0. It is desired to attach a host C to the network via a direct connection to B (see Figure 4.58). Explain how to do this with subnets; give sample subnet assignments. Assume that an additional network address is not available. What does this do to the size of the Ethernet LAN? 42 An alternative method for connecting host C in Exercise 41 is to use proxy ARP and routing: B agrees to route traffic to and from C and also answers ARP queries for C received over the Ethernet. (a) Give all packets sent, with physical addresses, as A uses ARP to locate and then send one packet to C. (b) Give B’s routing table. What peculiarity must it contain? 43 Propose a plausible addressing plan for IPv6 that runs out of bits. Specifically, provide a diagram such as Figure 4.33, perhaps with additional ID fields, that adds up to more than 128 bits, together with plausible justifications for the size of each field. You may assume fields are divided on byte boundaries and that Exercises 367 NetMaskLength NextHop C4.50.0.0/12 A C4.5E.10.0/20 B C4.60.0.0/12 C C4.68.0.0/14 D 80.0.0.0/1 E 40.0.0.0/2 F 00.0.0.0/2 G Table 4.16 Routing table for Exercise 45. the InterfaceID is 64 bits. Hint: Consider fields that would approach maximum allocation only under unusual circumstances. Can you do this if the InterfaceID is 48 bits? 44 Suppose two subnets share the same physical LAN; hosts on each subnet will see the other subnet’s broadcast packets. (a) How will DHCP fare if two servers, one for each subnet, coexist on the shared LAN? What problems might [do!] arise? (b) Will ARP be affected by such sharing? 45 Table 4.16 is a routing table using CIDR. Address bytes are in hexadecimal. The notation “/12” in C4.50.0.0/12 denotes a netmask with 12 leading 1 bits, that is, FF.F0.0.0. Note that the last three entries cover every address and thus serve in lieu of a default route. State to what next hop the following will be delivered. (a) C4.5E.13.87 (b) C4.5E.22.09 (c) C3.41.80.02 (d) 5E.43.91.12 (e) C4.6D.31.2E (f) C4.6B.31.2E 368 4 Internetworking NetMaskLength NextHop C4.5E.2.0/23 A C4.5E.4.0/22 B C4.5E.C0.0/19 C C4.5E.40.0/18 D C4.4C.0.0/14 E C0.0.0.0/2 F 80.0.0.0/1 G Table 4.17 Routing table for Exercise 46. 46 Table 4.17 is a routing table using CIDR. Address bytes are in hexadecimal. The notation “/12” in C4.50.0.0/12 denotes a netmask with 12 leading 1 bits, that is, FF.F0.0.0. State to what next hop the following will be delivered: (a) C4.4B.31.2E (b) C4.5E.05.09 (c) C4.4D.31.2E (d) C4.5E.03.87 (e) C4.5E.7F.12 (f) C4.5E.D1.02 47 Suppose P, Q, and R are network service providers, with respective CIDR address allocations (using the notation of Exercise 45) C1.0.0.0/8, C2.0.0.0/8, and C3.0.0.0/8. Each provider’s customers initially receive address allocations that are a subset of the provider’s. P has the following customers: PA, with allocation C1.A3.0.0/16, and PB, with allocation C1.B0.0.0/12. Q has the following customers: QA, with allocation C2.0A.10.0/20, and QB, with allocation C2.0B.0.0/16. Assume there are no other providers or customers. Exercises 369 (a) Give routing tables for P, Q, and R, assuming each provider connects to both of the others. (b) Now assume P is connected to Q and Q is connected to R, but P and R are not directly connected. Give tables for P and R. (c) Suppose customer PA acquires a direct link to Q, and QA acquires a direct link to P, in addition to existing links. Give tables for P and Q, ignoring R. 48 In the previous problem, assume each provider connects to both others. Suppose customer PA switches to provider Q and customer QB switches to provider R. Use the CIDR longest match rule to give routing tables for all three providers that allow PA and QB to switch without renumbering. 49 Suppose most of the Internet uses some form of geographical addressing, but that a large international organization has a single IP network address and routes its internal traffic over its own links. (a) Explain the routing inefficiency for the organization’s inbound traffic inherent in this situation. (b) Explain how the organization might solve this problem for outbound traffic. (c) For your method above to work for inbound traffic, what would have to happen? (d) Suppose the large organization now changes its addressing to separate geo- graphical addresses for each office. What will its internal routing structure have to look like if internal traffic is still to be routed internally? 50 The telephone system uses geographical addressing. Why do you think this wasn’t adopted as a matter of course by the Internet? 51 Suppose a site A is multihomed, in that it has two Internet connections from two different providers, P and Q. Provider-based addressing as in Exercise 47 is used, and A takes its address assignment from P. Q has a CIDR longest match routing entry for A. (a) Describe what inbound traffic might flow on the A–Q connection. Consider cases where Q does and does not advertise A to the world using BGP. (b) What is the minimum advertising of its route to A that Q must do in order for all inbound traffic to reach A via Q if the P–A link breaks? (c) What problems must be overcome if A is to use both links for its outbound traffic? 370 4 Internetworking 52 An ISP with a class B address is working with a new company to allocate it a portion of address space based on CIDR. The new company needs IP addresses for machines in three divisions of its corporate network: Engineering, Marketing, and Sales. These divisions plan to grow as follows: Engineering has 5 machines as of the start of year 1 and intends to add 1 machine every week; Marketing will never need more than 16 machines; and Sales needs 1 machine for every two clients. As of the start of year 1, the company has no clients, but the sales model indicates that by the start of year 2, the company will have six clients and each week thereafter gets one new client with probability 60%, loses one client with probability 20%, or maintains the same number with probability 20%. (a) What address range would be required to support the company’s growth plans for at least seven years if marketing uses all 16 of its addresses and the sales and engineering plans behave as expected? (b) How long would this address assignment last? At the time when the company runs out of address space, how would the addresses be assigned to the three groups? (c) If CIDR addressing were not available for the seven-year plan, what options would the new company have in terms of getting address space? 53 Propose a lookup algorithm for a CIDR fowarding table that does not require a linear search of the entire table to find the longest match. 54 Suppose a network N within a larger organization A acquires its own direct con- nection to an Internet service provider, in addition to an existing connection via A. Let R1 be the router connecting N to its own provider, and let R2 be the router connecting N to the rest of A. (a) Assuming N remains a subnet of A, how should R1 and R2 be configured? What limitations would still exist with N’s use of its separate connection? Would A be prevented from using N’s connection? Specify your configuration in terms of what R1 and R2 should advertise, and with what paths. Assume a BGP-like mechanism is available. (b) Now suppose N gets its own network number; how does this change your answer in (a)? (c) Describe a router configuration that would allow A to use N’s link when its own link is down. Exercises 371 R1 R2 R3 R4 R5 R6 R7 D E Figure 4.59 Example internet for Exercise 55. 55 Consider the example internet shown in Figure 4.59, in which sources D and E send packets to multicast group G, whose members are shaded in gray. Show the shortest-path multicast trees for each source. 56 Consider the example internet shown in Figure 4.60 in which sources S1 and S2 send packets to multicast group G, whose members are shaded in gray. Show the shortest-path multicast trees for each source. 57 Suppose host A is sending to a multicast group; the recipients are leaf nodes of a tree rooted at A with depth N and with each nonleaf node having k children; there are thus kN recipients. (a) How many individual link transmissions are involved if A sends a multicast message to all recipients? 372 4 Internetworking R1 S1 S2 R2 R4 R5 R6 R7 R8 Figure 4.60 Example network for Exercise 56. (b) How many individual link transmissions are involved if A sends unicast mes- sages to each individual recipient? (c) Suppose A sends to all recipients, but some messages are lost and retransmis- sion is necessary. Unicast retransmissions to what fraction of the recipients is equivalent, in terms of individual link transmissions, to a multicast retrans- mission to all recipients? 58 Determine whether or not the following IPv6 address notations are correct. (a) ::0F53:6382:AB00:67DB:BB27:7332 (b) 7803:42F2:::88EC:D4BA:B75D:11CD (c) ::4BA8:95CC::DB97:4EAB (d) 74DC::02BA (e) ::00FF:128.112.92.116 Exercises 373 59 Determine if your site is connected to the MBone. If so, investigate and experiment with any MBone tools, such as sdr, vat, and vic. 60 MPLS labels are usually 20 bits long. Explain why this provides enough labels when MPLS is used for destination-based forwarding. 61 MPLS has sometimes been claimed to improve router performance. Explain why this might be true, and suggest reasons why in practice this may not be the case. 62 Assume that it takes 32 bits to carry each MPLS label that is added to a packet when the “shim” header of Figure 4.42(b) is used. (a) How many additional bytes are needed to tunnel a packet using the MPLS techniques described in Section 4.5.3? (b) How many additional bytes are needed, at a minimum, to tunnel a packet using an additional IP header as described in Section 4.1.8? (c) Calculate the efficiency of bandwidth usage for each of the two tunneling approaches when the average packet size is 300 bytes. Repeat for 64-byte packets. Bandwidth efficiency is defined as (payload bytes carried) ÷ (total bytes carried). 63 RFC 791 describes the Internet Protocol and includes two options for source routing. Describe three disadvantages of using IP source route options compared to using MPLS for explicit routing. (Hint: The IP header including options may be at most 15 words long.) End-to-End Protocols Victory is the beautiful, bright coloured flower. Transport is the stem without which it could never have blossomed. —Winston Churchill The previous three chapters have described various technologies that can be used to connect together a collection of computers: direct links (including LAN technologies like Ethernet and token ring), packet-switched networks (including cell-based networks like ATM), and internetworks. The next problem is to turn this host-to-host packet delivery service into a process-to-process communication PROBLEM Getting Processes to Communicate channel. This is the role played by the transport level of the network archi- tecture, which, because it supports communication between the end application programs, is sometimes called the end-to-end protocol. Two forces shape the end-to-end protocol. From above, the application-level processes that use its services have cer- tain requirements. The following list itemizes some of the common properties that a transport protocol can be expected to provide: ■ guarantees message delivery ■ delivers messages in the same order they are sent ■ delivers at most one copy of each message ■ supports arbitrarily large messages ■ supports synchronization between the sender and the receiver ■ allows the receiver to apply flow control to the sender ■ supports multiple application processes on each host 5Note that this list does not include all the functionality that application processes might want from the network. For example, it does not include security, which is typically provided by protocols that sit above the transport level. From below, the underlying network upon which the transport protocol operates has certain limitations in the level of service it can provide. Some of the more typical limitations of the network are that it may ■ drop messages ■ reorder messages ■ deliver duplicate copies of a given message ■ limit messages to some finite size ■ deliver messages after an arbitrarily long delay Such a network is said to provide a best-effort level of service, as exemplified by the Internet. The challenge, therefore, is to develop algorithms that turn the less-than-desirable properties of the underly- ing network into the high level of service required by ap- plication programs. Different transport protocols employ different combinations of these algorithms. This chapter looks at these algorithms in the context of three repre- sentative services—a simple asynchronous demultiplexing service, a reliable byte-stream service, and a request/reply service. In the case of the demultiplexing and byte-stream services, we use the Internet’s UDP and TCP protocols, respectively, to illustrate how these services are provided in practice. In the third case, we first give a collection of algorithms that implement the request/reply (plus other re- lated) services and then show how these algorithms can be combined to implement a Remote Procedure Call (RPC) protocol. This discussion is capped off with a description of two widely used RPC protocols—SunRPC and DCE- RPC—in terms of these component algorithms. Finally, the chapter concludes with a section that discusses the performance of the different transport protocols. 376 5 End-to-End Protocols 5.1 Simple Demultiplexer (UDP) The simplest possible transport protocol is one that extends the host-to-host delivery service of the underlying network into a process-to-process communication service. There are likely to be many processes running on any given host, so the protocol needs to add a level of demultiplexing, thereby allowing multiple application processes on each host to share the network. Aside from this requirement, the transport protocol adds no other functionality to the best-effort service provided by the underlying net- work. The Internet’s User Datagram Protocol (UDP) is an example of such a transport protocol. The only interesting issue in such a protocol is the form of the address used to identify the target process. Although it is possible for processes to directly identify each other with an OS-assigned process id (pid), such an approach is only practical in a closed distributed system in which a single OS runs on all hosts and assigns each process a unique id. A more common approach, and the one used by UDP, is for processes to indirectly identify each other using an abstract locator, often called a port or mailbox. The basic idea is for a source process to send a message to a port and for the destination process to receive the message from a port. The header for an end-to-end protocol that implements this demultiplexing func- tion typically contains an identifier (port) for both the sender (source) and the receiver (destination) of the message. For example, the UDP header is given in Figure 5.1. Notice that the UDP port field is only 16 bits long. This means that there are up to 64K possi- ble ports, clearly not enough to identify all the processes on all the hosts in the Internet. Fortunately, ports are not interpreted across the entire Internet, but only on a single host. That is, a process is really identified by a port on some particular host—a port, host pair. In fact, this pair constitutes the demultiplexing key for the UDP protocol. The next issue is how a process learns the port for the process to which it wants to send a message. Typically, a client process initiates a message exchange with a server SrcPort DstPort ChecksumLength Data 01631 Figure 5.1 Format for UDP header. 5.1 Simple Demultiplexer (UDP) 377 process. Once a client has contacted a server, the server knows the client’s port (it was contained in the message header) and can reply to it. The real problem, therefore, is how the client learns the server’s port in the first place. A common approach is for the server to accept messages at a well-known port. That is, each server receives its messages at some fixed port that is widely published, much like the emergency telephone service available at the well-known phone number 911. In the Internet, for example, the Domain Name Server (DNS) receives messages at well-known port 53 on each host, the mail service listens for messages at port 25, and the Unix talk program accepts messages at well-known port 517, and so on. This mapping is published periodically in an RFC and is available on most Unix systems in file /etc/services. Sometimes a well-known port is just the starting point for communication: The client and server use the well-known port to agree on some other port that they will use for subsequent communication, leaving the well-known port free for other clients. An alternative strategy is to generalize this idea, so that there is only a single well-known port—the one at which the “Port Mapper” service accepts messages. A client would send a message to the Port Mapper’s well-known port asking for the port it should use to talk to the “whatever” service, and the Port Mapper returns the appropriate port. This strategy makes it easy to change the port associated with different services over time, and for each host to use a different port for the same service. As just mentioned, a port is purely an abstraction. Exactly how it is implemented differs from system to system, or more precisely, from OS to OS. For example, the socket API described in Chapter 1 is an implementation of ports. Typically, a port is implemented by a message queue, as illustrated in Figure 5.2. When a message arrives, the protocol (e.g., UDP) appends the message to the end of the queue. Should the queue be full, the message is discarded. There is no flow-control mechanism that tells the sender to slow down. When an application process wants to receive a message, one is removed from the front of the queue. If the queue is empty, the process blocks until a message becomes available. Finally, although UDP does not implement flow control or reliable/ordered deliv- ery, it does a little more work than to simply demultiplex messages to some application process—it also ensures the correctness of the message by the use of a checksum. (The UDP checksum is optional in the current Internet, but it will become mandatory with IPv6.) UDP computes its checksum over the UDP header, the contents of the message body, and something called the pseudoheader. The pseudoheader consists of three fields from the IP header—protocol number, source IP address, and destination IP address— plus the UDP length field. (Yes, the UDP length field is included twice in the checksum calculation.) UDP uses the same checksum algorithm as IP, as defined in Section 2.4.2. The motivation behind having the pseudoheader is to verify that this message has been 378 5 End-to-End Protocols Application process Application process Application process UDP Packets arrive Ports Queues Packets demultiplexed Figure 5.2 UDP message queue. delivered between the correct two endpoints. For example, if the destination IP address was modified while the packet was in transit, causing the packet to be misdelivered, this fact would be detected by the UDP checksum. 5.2 Reliable Byte Stream (TCP) In contrast to a simple demultiplexing protocol like UDP, a more sophisticated trans- port protocol is one that offers a reliable, connection-oriented, byte-stream service. Such a service has proven useful to a wide assortment of applications because it frees the application from having to worry about missing or reordered data. The Internet’s Transmission Control Protocol (TCP) is probably the most widely used protocol of this type; it is also the most carefully tuned. It is for these two reasons that this section studies TCP in detail, although we identify and discuss alternative design choices at the end of the section. In terms of the properties of transport protocols given in the problem statement at the start of this chapter, TCP guarantees the reliable, in-order delivery of a stream of bytes. It is a full-duplex protocol, meaning that each TCP connection supports a 5.2 Reliable Byte Stream (TCP) 379 pair of byte streams, one flowing in each direction. It also includes a flow-control mechanism for each of these byte streams that allows the receiver to limit how much data the sender can transmit at a given time. Finally, like UDP, TCP supports a de- multiplexing mechanism that allows multiple application programs on any given host to simultaneously carry on a conversation with their peers. In addition to the above features, TCP also implements a highly tuned congestion-control mechanism. The idea of this mechanism is to throttle how fast TCP sends data, not for the sake of keeping the sender from overrunning the receiver, but to keep the sender from overloading the network. A description of TCP’s congestion-control mechanism is postponed until Chapter 6, where we discuss it in the larger context of how network resources are fairly allocated.◮ Since many people confuse congestion control and flow control, we restate the difference. Flow control involves preventing senders from overrunning the capacity of receivers. Congestion control involves preventing too much data from being injected into the network, thereby causing switches or links to become overloaded. Thus, flow control is an end-to-end issue, while congestion control is concerned with how hosts and networks interact. 5.2.1 End-to-End Issues At the heart of TCP is the sliding window algorithm. Even though this is the same basic algorithm we saw in Section 2.5.2, because TCP runs over the Internet rather than a point-to-point link, there are many important differences. This subsection identifies these differences and explains how they complicate TCP. The following subsections then describe how TCP addresses these and other complications. First, whereas the sliding window algorithm presented in Section 2.5.2 runs over a single physical link that always connects the same two computers, TCP supports logical connections between processes that are running on any two computers in the Internet. This means that TCP needs an explicit connection establishment phase during which the two sides of the connection agree to exchange data with each other. This difference is analogous to having to dial up the other party, rather than having a dedicated phone line. TCP also has an explicit connection teardown phase. One of the things that happens during connection establishment is that the two parties establish some shared state to enable the sliding window algorithm to begin. Connection teardown is needed so each host knows it is OK to free this state. Second, whereas a single physical link that always connects the same two com- puters has a fixed RTT, TCP connections are likely to have widely different round-trip times. For example, a TCP connection between a host in San Francisco and a host in Boston, which are separated by several thousand kilometers, might have an RTT 380 5 End-to-End Protocols of 100 ms, while a TCP connection between two hosts in the same room, only a few meters apart, might have an RTT of only 1 ms. The same TCP protocol must be able to support both of these connections. To make matters worse, the TCP connection between hosts in San Francisco and Boston might have an RTT of 100 ms at 3 a.m., but an RTT of 500 ms at 3 p.m. Variations in the RTT are even possible during a single TCP connection that lasts only a few minutes. What this means to the sliding window algorithm is that the timeout mechanism that triggers retransmissions must be adaptive. (Certainly, the timeout for a point-to-point link must be a settable parameter, but it is not necessary to adapt this timer for a particular pair of nodes.) A third difference is that packets may be reordered as they cross the Internet, but this is not possible on a point-to-point link where the first packet put into one end of the link must be the first to appear at the other end. Packets that are slightly out of order do not cause a problem since the sliding window algorithm can reorder packets correctly using the sequence number. The real issue is how far out-of-order packets can get, or said another way, how late a packet can arrive at the destination. In the worst case, a packet can be delayed in the Internet until IP’s time to live (TTL) field expires, at which time the packet is discarded (and hence there is no danger of it arriving late). Knowing that IP throws packets away after their TTL expires, TCP assumes that each packet has a maximum lifetime. The exact lifetime, known as the maximum segment lifetime (MSL), is an engineering choice. The current recommended setting is 120 seconds. Keep in mind that IP does not directly enforce this 120-second value; it is simply a conservative estimate that TCP makes of how long a packet might live in the Internet. The implication is significant—TCP has to be prepared for very old packets to suddenly show up at the receiver, potentially confusing the sliding window algorithm. Fourth, the computers connected to a point-to-point link are generally engineered to support the link. For example, if a link’s delay × bandwidth product is computed to be 8 KB—meaning that a window size is selected to allow up to 8 KB of data to be unacknowledged at a given time—then it is likely that the computers at either end of the link have the ability to buffer up to 8 KB of data. Designing the system otherwise would be silly. On the other hand, almost any kind of computer can be connected to the Internet, making the amount of resources dedicated to any one TCP connection highly variable, especially considering that any one host can potentially support hundreds of TCP connections at the same time. This means that TCP must include a mechanism that each side uses to “learn” what resources (e.g., how much buffer space) the other side is able to apply to the connection. This is the flow-control issue. Fifth, because the transmitting side of a directly connected link cannot send any faster than the bandwidth of the link allows, and only one host is pumping data into the link, it is not possible to unknowingly congest the link. Said another way, the load 5.2 Reliable Byte Stream (TCP) 381 on the link is visible in the form of a queue of packets at the sender. In contrast, the sending side of a TCP connection has no idea what links will be traversed to reach the destination. For example, the sending machine might be directly connected to a relatively fast Ethernet—and so, capable of sending data at a rate of 100 Mbps—but somewhere out in the middle of the network, a 1.5-Mbps T1 link must be traversed. And to make matters worse, data being generated by many different sources might be trying to traverse this same slow link. This leads to the problem of network congestion. Discussion of this topic is delayed until Chapter 6. We conclude this discussion of end-to-end issues by comparing TCP’s approach to providing a reliable/ordered delivery service with the approach used by X.25 networks. In TCP, the underlying IP network is assumed to be unreliable and to deliver messages out of order; TCP uses the sliding window algorithm on an end-to-end basis to provide reliable/ordered delivery. In contrast, X.25 networks use the sliding window protocol within the network, on a hop-by-hop basis. The assumption behind this approach is that if messages are delivered reliably and in order between each pair of nodes along the path between the source host and the destination host, then the end-to-end service also guarantees reliable/ordered delivery. The problem with this latter approach is that a sequence of hop-by-hop guaran- tees does not necessarily add up to an end-to-end guarantee. First, if a heterogeneous link (say, an Ethernet) is added to one end of the path, then there is no guarantee that this hop will preserve the same service as the other hops. Second, just because the sliding window protocol guarantees that messages are delivered correctly from node A to node B, and then from node B to node C, it does not guarantee that node B behaves perfectly. For example, network nodes have been known to introduce errors into messages while transferring them from an input buffer to an output buffer. They have also been known to accidentally reorder messages. As a consequence of these small windows of vulnerability, it is still necessary to provide true end-to-end checks to guarantee reliable/ordered service, even though the lower levels of the system also implement that functionality.◮ This discussion serves to illustrate one of the most important principles in system design—the end-to-end argument. In a nutshell, the end-to-end argument says that a function (in our example, providing reliable/ordered delivery) should not be provided in the lower levels of the system unless it can be completely and correctly implemented at that level. Therefore, this rule argues in favor of the TCP/IP approach. This rule is not absolute, however. It does allow for functions to be incompletely provided at a low level as a performance optimization. This is why it is perfectly consistent with the end-to-end argument to perform error detection (e.g., CRC) on a hop-by-hop basis; detecting and retransmitting a single corrupt packet across one hop is preferable to having to retransmit an entire file end-to-end. 382 5 End-to-End Protocols Application process Write bytes TCP Send buffer Segment Segment Segment Transmit segments Application process Read bytes TCP Receive buffer … … … Figure 5.3 How TCP manages a byte stream. 5.2.2 Segment Format TCP is a byte-oriented protocol, which means that the sender writes bytes into a TCP connection and the receiver reads bytes out of the TCP connection. Although “byte stream” describes the service TCP offers to application processes, TCP does not, itself, transmit individual bytes over the Internet. Instead, TCP on the source host buffers enough bytes from the sending process to fill a reasonably sized packet and then sends this packet to its peer on the destination host. TCP on the destination host then empties the contents of the packet into a receive buffer, and the receiving process reads from this buffer at its leisure. This situation is illustrated in Figure 5.3, which, for simplicity, shows data flowing in only one direction. Remember that, in general, a single TCP connection supports byte streams flowing in both directions. The packets exchanged between TCP peers in Figure 5.3 are called segments, since each one carries a segment of the byte stream. Each TCP segment contains the header schematically depicted in Figure 5.4. The relevance of most of these fields will become apparent throughout this section. For now, we simply introduce them. The SrcPort and DstPort fields identify the source and destination ports, respec- tively, just as in UDP. These two fields, plus the source and destination IP addresses, combine to uniquely identify each TCP connection. That is, TCP’s demux key is given by the 4-tuple  SrcPort, SrcIPAddr, DstPort, DstIPAddr  Note that because TCP connections come and go, it is possible for a connection be- tween a particular pair of ports to be established, used to send and receive data, and closed, and then at a later time for the same pair of ports to be involved in a second 5.2 Reliable Byte Stream (TCP) 383 Options (variable) Data Checksum SrcPort DstPort HdrLen 0 Flags UrgPtr AdvertisedWindow SequenceNum Acknowledgment 04 10 16 31 Figure 5.4 TCP header format. Sender Data (SequenceNum) Acknowledgment + AdvertisedWindow Receiver Figure 5.5 Simplified illustration (showing only one direction) of the TCP process, with data flow in one direction and ACKs in the other. connection. We sometimes refer to this situation as two different incarnations of the same connection. The Acknowledgment, SequenceNum, and AdvertisedWindow fields are all in- volved in TCP’s sliding window algorithm. Because TCP is a byte-oriented protocol, each byte of data has a sequence number; the SequenceNum field contains the sequence number for the first byte of data carried in that segment. The Acknowledgment and AdvertisedWindow fields carry information about the flow of data going in the other direction. To simplify our discussion, we ignore the fact that data can flow in both directions, and we concentrate on data that has a particular SequenceNum flowing in one direction and Acknowledgment and AdvertisedWindow values flowing in the opposite direction, as illustrated in Figure 5.5. The use of these three fields is described more fully in Section 5.2.4. The 6-bit Flags field is used to relay control information between TCP peers. The possible flags include SYN, FIN, RESET, PUSH, URG, and ACK. The SYN and FIN flags 384 5 End-to-End Protocols are used when establishing and terminating a TCP connection, respectively. Their use is described in Section 5.2.3. The ACK flag is set any time the Acknowledgment field is valid, implying that the receiver should pay attention to it. The URG flag signifies that this segment contains urgent data. When this flag is set, the UrgPtr field indicates where the nonurgent data contained in this segment begins. The urgent data is contained at the front of the segment body, up to and including a value of UrgPtr bytes into the segment. The PUSH flag signifies that the sender invoked the push operation, which indicates to the receiving side of TCP that it should notify the receiving process of this fact. We discuss these last two features more in Section 5.2.7. Finally, the RESET flag signifies that the receiver has become confused—for example, because it received a segment it did not expect to receive—and so wants to abort the connection. Finally, the Checksum field is used in exactly the same way as for UDP—it is computed over the TCP header, the TCP data, and the pseudoheader, which is made up of the source address, destination address, and length fields from the IP header. The checksum is required for TCP in both IPv4 and IPv6. Also, since the TCP header is of variable length (options can be attached after the mandatory fields), a HdrLen field is included that gives the length of the header in 32-bit words. This field is also known as the Offset field, since it measures the offset from the start of the packet to the start of the data. 5.2.3 Connection Establishment and Termination A TCP connection begins with a client (caller) doing an active open to a server (callee). Assuming that the server had earlier done a passive open, the two sides engage in an exchange of messages to establish the connection. (Recall from Chapter 1 that a party wanting to initiate a connection performs an active open, while a party will- ing to accept a connection does a passive open.) Only after this connection estab- lishment phase is over do the two sides begin sending data. Likewise, as soon as a participant is done sending data, it closes one direction of the connection, which causes TCP to initiate a round of connection termination messages. Notice that while connection setup is an asymmetric activity (one side does a passive open and the other side does an active open), connection teardown is symmetric (each side has to close the connection independently).1 Therefore, it is possible for one side to have done a close, meaning that it can no longer send data, but for the other side to keep the other half of the bidirectional connection open and to continue sending data. 1To be more precise, connection setup can be symmetric, with both sides trying to open the connection at the same time, but the common case is for one side to do an active open and the other side to do a passive open. 5.2 Reliable Byte Stream (TCP) 385 Active participant (client) Passive participant (server) SYN, SequenceNum = x SYN + ACK, SequenceNum = y, ACK, Acknowledgmen t = y + 1 Acknowledgment = x + 1 Figure 5.6 Timeline for three-way handshake algorithm. Three-Way Handshake The algorithm used by TCP to establish and terminate a connection is called a three- way handshake. We first describe the basic algorithm and then show how it is used by TCP. The three-way handshake involves the exchange of three messages between the client and the server, as illustrated by the timeline given in Figure 5.6. The idea is that two parties want to agree on a set of parameters, which, in the case of opening a TCP connection, are the starting sequence numbers the two sides plan to use for their respective byte streams. In general, the parameters might be any facts that each side wants the other to know about. First, the client (the active participant) sends a segment to the server (the passive participant) stating the initial sequence number it plans to use (Flags = SYN, SequenceNum = x). The server then responds with a single segment that both acknowledges the client’s sequence number (Flags =ACK, Ack = x + 1) and states its own beginning sequence number (Flags = SYN, SequenceNum = y). That is, both the SYN and ACK bits are set in the Flags field of this second message. Finally, the client responds with a third segment that acknowledges the server’s sequence number (Flags = ACK, Ack = y + 1). The reason that each side acknowledges a sequence number that is one larger than the one sent is that the Acknowledgment field actually identifies the “next sequence number expected,” thereby implicitly acknowledging all earlier sequence numbers. Although not shown in this timeline, a timer is scheduled for each of the first two segments, and if the expected response is not received, the segment is retransmitted. You may be asking yourself why the client and server have to exchange starting sequence numbers with each other at connection setup time. It would be simpler if each side simply started at some “well-known” sequence number, such as 0. In fact, 386 5 End-to-End Protocols the TCP specification requires that each side of a connection select an initial starting sequence number at random. The reason for this is to protect against two incarnations of the same connection reusing the same sequence numbers too soon, that is, while there is still a chance that a segment from an earlier incarnation of a connection might interfere with a later incarnation of the connection. State Transition Diagram TCP is complex enough that its specification includes a state transition diagram. A copy of this diagram is given in Figure 5.7. This diagram shows only the states in- volved in opening a connection (everything above ESTABLISHED) and in closing a connection (everything below ESTABLISHED). Everything that goes on while a con- nection is open—that is, the operation of the sliding window algorithm—is hidden in the ESTABLISHED state. CLOSED LISTEN SYN_RCVD SYN_SENT ESTABLISHED CLOSE_WAIT LAST_ACKCLOSING TIME_WAIT FIN_WAIT_2 FIN_WAIT_1 Passive open Close Send/SYN SYN/SYN + ACK SYN + ACK/ACK SYN/SYN + ACK ACK Close/FIN FIN/ACKClose/FIN FIN/ACKACK + FIN/ACK Timeout after two segment lifetimesFIN/ACK ACK ACK ACK Close/FIN Close CLOSED Active open/SYN Figure 5.7 TCP state transition diagram. 5.2 Reliable Byte Stream (TCP) 387 TCP’s state transition diagram is fairly easy to understand. Each circle denotes a state that one end of a TCP connection can find itself in. All connections start in the CLOSED state. As the connection progresses, the connection moves from state to state according to the arcs. Each arc is labelled with a tag of the form event/action. Thus, if a connection is in the LISTEN state and a SYN segment arrives (i.e., a segment with the SYN flag set), the connection makes a transition to the SYN RCVD state and takes the action of replying with an ACK + SYN segment. Notice that two kinds of events trigger a state transition: (1) a segment arrives from the peer (e.g., the event on the arc from LISTEN to SYN RCVD), or (2) the local application process invokes an operation on TCP (e.g., the active open event on the arc from CLOSE to SYN SENT). In other words, TCP’s state transition diagram effectively defines the semantics of both its peer-to-peer interface and its service interface, as defined in Section 1.3.1. The syntax of these two interfaces is given by the segment format (as illustrated in Figure 5.4) and by some application programming interface (an example of which is given in Section 1.4.1), respectively. Now let’s trace the typical transitions taken through the diagram in Figure 5.7. Keep in mind that at each end of the connection, TCP makes different transitions from state to state. When opening a connection, the server first invokes a passive open operation on TCP, which causes TCP to move to the LISTEN state. At some later time, the client does an active open, which causes its end of the connection to send a SYN segment to the server and to move to the SYN SENT state. When the SYN segment arrives at the server, it moves to the SYN RCVD state and responds with a SYN+ACK segment. The arrival of this segment causes the client to move to the ESTABLISHED state and to send an ACK back to the server. When this ACK arrives, the server finally moves to the ESTABLISHED state. In other words, we have just traced the three-way handshake. There are three things to notice about the connection establishment half of the state transition diagram. First, if the client’s ACK to the server is lost, corresponding to the third leg of the three-way handshake, then the connection still functions correctly. This is because the client side is already in the ESTABLISHED state, so the local application process can start sending data to the other end. Each of these data segments will have the ACK flag set, and the correct value in the Acknowledgment field, so the server will move to the ESTABLISHED state when the first data segment arrives. This is actually an important point about TCP—every segment reports what sequence number the sender is expecting to see next, even if this repeats the same sequence number contained in one or more previous segments. The second thing to notice about the state transition diagram is that there is a funny transition out of the LISTEN state whenever the local process invokes a send operation on TCP. That is, it is possible for a passive participant to identify both ends 388 5 End-to-End Protocols of the connection (i.e., itself and the remote participant that it is willing to have connect to it), and then to change its mind about waiting for the other side and instead actively establish the connection. To the best of our knowledge, this is a feature of TCP that no application process actually takes advantage of. The final thing to notice about the diagram is the arcs that are not shown. Specif- ically, most of the states that involve sending a segment to the other side also schedule a timeout that eventually causes the segment to be resent if the expected response does not happen. These retransmissions are not depicted in the state transition diagram. If after several tries the expected response does not arrive, TCP gives up and returns to the CLOSED state. Turning our attention now to the process of terminating a connection, the im- portant thing to keep in mind is that the application process on both sides of the connection must independently close its half of the connection. If only one side closes the connection, then this means it has no more data to send, but it is still available to receive data from the other side. This complicates the state transition diagram be- cause it must account for the possibility that the two sides invoke the close operator at the same time, as well as the possibility that first one side invokes close and then, at some later time, the other side invokes close. Thus, on any one side there are three combinations of transitions that get a connection from the ESTABLISHED state to the CLOSED state: ■ This side closes first: ESTABLISHED → FIN WAIT 1 → FIN WAIT 2 → TIME WAIT →CLOSED. ■ The other side closes first: ESTABLISHED → CLOSE WAIT → LAST ACK → CLOSED. ■ Both sides close at the same time: ESTABLISHED → FIN WAIT 1 → CLOSING → TIME WAIT →CLOSED. There is actually a fourth, although rare, sequence of transitions that leads to the CLOSED state; it follows the arc from FIN WAIT 1 to TIME WAIT. We leave it as an exercise for you to figure out what combination of circumstances leads to this fourth possibility. The main thing to recognize about connection teardown is that a connection in the TIME WAIT state cannot move to the CLOSED state until it has waited for two times the maximum amount of time an IP datagram might live in the Internet (i.e., 120 seconds). The reason for this is that while the local side of the connection has sent an ACK in response to the other side’s FIN segment, it does not know that the ACK was successfully delivered. As a consequence, the other side might retransmit its 5.2 Reliable Byte Stream (TCP) 389 FIN segment, and this second FIN segment might be delayed in the network. If the connection were allowed to move directly to the CLOSED state, then another pair of application processes might come along and open the same connection (i.e., use the same pair of port numbers), and the delayed FIN segment from the earlier incarnation of the connection would immediately initiate the termination of the later incarnation of that connection. 5.2.4 Sliding Window Revisited We are now ready to discuss TCP’s variant of the sliding window algorithm, which serves several purposes: (1) it guarantees the reliable delivery of data, (2) it ensures that data is delivered in order, and (3) it enforces flow control between the sender and the receiver. TCP’s use of the sliding window algorithm is the same as we saw in Section 2.5.2 in the case of the first two of these three functions. Where TCP differs from the earlier algorithm is that it folds the flow-control function in as well. In particular, rather than having a fixed-size sliding window, the receiver advertises a window size to the sender. This is done using the AdvertisedWindow field in the TCP header. The sender is then limited to having no more than a value of AdvertisedWindow bytes of unacknowledged data at any given time. The receiver selects a suitable value for AdvertisedWindow based on the amount of memory allocated to the connection for the purpose of buffering data. The idea is to keep the sender from overrunning the receiver’s buffer. We discuss this at greater length below. Reliable and Ordered Delivery To see how the sending and receiving sides of TCP interact with each other to im- plement reliable and ordered delivery, consider the situation illustrated in Figure 5.8. TCP on the sending side maintains a send buffer. This buffer is used to store data Sending application LastByteWritten TCP LastByteSentLastByteAcked Receiving application LastByteRead TCP LastByteRcvdNextByteExpected (a) (b) Figure 5.8 Relationship between TCP send buffer (a) and receive buffer (b). 390 5 End-to-End Protocols that has been sent but not yet acknowledged, as well as data that has been written by the sending application, but not transmitted. On the receiving side, TCP maintains a receive buffer. This buffer holds data that arrives out of order, as well as data that is in the correct order (i.e., there are no missing bytes earlier in the stream) but that the application process has not yet had the chance to read. To make the following discussion simpler to follow, we initially ignore the fact that both the buffers and the sequence numbers are of some finite size and hence will eventually wrap around. Also, we do not distinguish between a pointer into a buffer where a particular byte of data is stored and the sequence number for that byte. Looking first at the sending side, three pointers are maintained into the send buf- fer, each with an obvious meaning: LastByteAcked, LastByteSent, and LastByteWritten. Clearly, LastByteAcked ≤ LastByteSent since the receiver cannot have acknowledged a byte that has not yet been sent, and LastByteSent ≤ LastByteWritten since TCP cannot send a byte that the application process has not yet written. Also note that none of the bytes to the left of LastByteAcked need to be saved in the buffer because they have already been acknowledged, and none of the bytes to the right of LastByteWritten need to be buffered because they have not yet been generated. A similar set of pointers (sequence numbers) are maintained on the receiving side: LastByteRead, NextByteExpected, and LastByteRcvd. The inequalities are a little less in- tuitive, however, because of the problem of out-of-order delivery. The first relationship LastByteRead < NextByteExpected is true because a byte cannot be read by the application until it is received and all pre- ceding bytes have also been received. NextByteExpected points to the byte immediately after the latest byte to meet this criterion. Second, NextByteExpected ≤ LastByteRcvd + 1 since, if data has arrived in order, NextByteExpected points to the byte after LastByte- Rcvd, whereas if data has arrived out of order, NextByteExpected points to the start of the first gap in the data, as in Figure 5.8. Note that bytes to the left of LastByteRead need not be buffered because they have already been read by the local application process, and bytes to the right of LastByteRcvd need not be buffered because they have not yet arrived. Flow Control Most of the above discussion is similar to that found in Section 2.5.2; the only real difference is that this time we elaborated on the fact that the sending and receiving ap- plication processes are filling and emptying their local buffer, respectively. (The earlier 5.2 Reliable Byte Stream (TCP) 391 discussion glossed over the fact that data arriving from an upstream node was filling the send buffer, and data being transmitted to a downstream node was emptying the receive buffer.) You should make sure you understand this much before proceeding because now comes the point where the two algorithms differ more significantly. In what follows, we reintroduce the fact that both buffers are of some finite size, denoted MaxSendBuffer and MaxRcvBuffer, although we don’t worry about the details of how they are implemented. In other words, we are only interested in the number of bytes being buffered, not in where those bytes are actually stored. Recall that in a sliding window protocol, the size of the window sets the amount of data that can be sent without waiting for acknowledgment from the receiver. Thus, the receiver throttles the sender by advertising a window that is no larger than the amount of data that it can buffer. Observe that TCP on the receive side must keep LastByteRcvd − LastByteRead ≤ MaxRcvBuffer to avoid overflowing its buffer. It therefore advertises a window size of AdvertisedWindow = MaxRcvBuffer − ((NextByteExpected − 1) − LastByteRead) which represents the amount of free space remaining in its buffer. As data arrives, the receiver acknowledges it as long as all the preceding bytes have also arrived. In addition, LastByteRcvd moves to the right (is incremented), meaning that the advertised window potentially shrinks. Whether or not it shrinks depends on how fast the local application process is consuming data. If the local process is reading data just as fast as it arrives (causing LastByteRead to be incremented at the same rate as LastByteRcvd), then the advertised window stays open (i.e., AdvertisedWindow = MaxRcvBuffer). If, however, the receiving process falls behind, perhaps because it performs a very expensive operation on each byte of data that it reads, then the advertised window grows smaller with every segment that arrives, until it eventually goes to 0. TCP on the send side must then adhere to the advertised window it gets from the receiver. This means that at any given time, it must ensure that LastByteSent − LastByteAcked ≤ AdvertisedWindow Said another way, the sender computes an effective window that limits how much data it can send: EffectiveWindow = AdvertisedWindow − (LastByteSent − LastByteAcked) Clearly, EffectiveWindow must be greater than 0 before the source can send more data. It is possible, therefore, that a segment arrives acknowledging x bytes, thereby allowing the sender to increment LastByteAcked by x, but because the receiving process was not reading any data, the advertised window is now x bytes smaller than the time before. 392 5 End-to-End Protocols In such a situation, the sender would be able to free buffer space, but not to send any more data. All the while this is going on, the send side must also make sure that the local application process does not overflow the send buffer, that is, that LastByteWritten − LastByteAcked ≤ MaxSendBuffer If the sending process tries to write y bytes to TCP, but (LastByteWritten − LastByteAcked) + y > MaxSendBuffer then TCP blocks the sending process and does not allow it to generate more data. It is now possible to understand how a slow receiving process ultimately stops a fast sending process. First, the receive buffer fills up, which means the advertised window shrinks to 0. An advertised window of 0 means that the sending side cannot transmit any data, even though data it has previously sent has been successfully ac- knowledged. Finally, not being able to transmit any data means that the send buffer fills up, which ultimately causes TCP to block the sending process. As soon as the receiving process starts to read data again, the receive-side TCP is able to open its win- dow back up, which allows the send-side TCP to transmit data out of its buffer. When this data is eventually acknowledged, LastByteAcked is incremented, the buffer space holding this acknowledged data becomes free, and the sending process is unblocked and allowed to proceed. There is only one remaining detail that must be resolved—how does the sending side know that the advertised window is no longer 0? As mentioned above, TCP always sends a segment in response to a received data segment, and this response contains the latest values for the Acknowledge and AdvertisedWindow fields, even if these values have not changed since the last time they were sent. The problem is this. Once the receive side has advertised a window size of 0, the sender is not permitted to send any more data, which means it has no way to discover that the advertised window is no longer 0 at some time in the future. TCP on the receive side does not sponta- neously send nondata segments; it only sends them in response to an arriving data segment. TCP deals with this situation as follows. Whenever the other side advertises a window size of 0, the sending side persists in sending a segment with 1 byte of data every so often. It knows that this data will probably not be accepted, but it tries anyway, because each of these 1-byte segments triggers a response that contains the current advertised window. Eventually, one of these 1-byte probes triggers a response that reports a nonzero advertised window.◮ Note that the reason the sending side periodically sends this probe segment is that TCP is designed to make the receive side as simple as possible—it simply responds 5.2 Reliable Byte Stream (TCP) 393 to segments from the sender, and it never initiates any activity on its own. This is an example of a well-recognized (although not universally applied) protocol design rule, which, for lack of a better name, we call the smart sender/dumb receiver rule. Recall that we saw another example of this rule when we discussed the use of NAKs in Section 2.5.2. Protecting against Wraparound This subsection and the next consider the size of the SequenceNum and AdvertisedWin- dow fields and the implications of their sizes on TCP’s correctness and performance. TCP’s SequenceNum field is 32 bits long, and its AdvertisedWindow field is 16 bits long, meaning that TCP has easily satisfied the requirement of the sliding window algo- rithm that the sequence number space be twice as big as the window size: 232 ≫ 2×216. However, this requirement is not the interesting thing about these two fields. Consider each field in turn. The relevance of the 32-bit sequence number space is that the sequence number used on a given connection might wrap around—a byte with sequence number x could be sent at one time, and then at a later time a second byte with the same sequence number x might be sent. Once again, we assume that packets cannot survive in the Internet for longer than the recommended MSL. Thus, we currently need to make sure that the sequence number does not wrap around within a 120-second period of time. Whether or not this happens depends on how fast data can be transmitted over the Internet, that is, how fast the 32-bit sequence number space can be consumed. (This discussion assumes that we are trying to consume the sequence number space as fast as possible, but of course we will be if we are doing our job of keeping the pipe full.) Table 5.1 shows how long it takes for the sequence number to wrap around on networks with various bandwidths. As you can see, the 32-bit sequence number space is adequate for today’s net- works, but given that OC-48 links currently exist in the Internet backbone, it won’t be long until individual TCP connections want to run at 622-Mbps speeds or higher. Fortunately, the IETF has already worked out an extension to TCP that effectively extends the sequence number space to protect against the sequence number wrapping around. This and related extensions are described in Section 5.2.8. Keeping the Pipe Full The relevance of the 16-bit AdvertisedWindow field is that it must be big enough to allow the sender to keep the pipe full. Clearly, the receiver is free not to open the window as large as the AdvertisedWindow field allows; we are interested in the situation in which the receiver has enough buffer space to handle as much data as the largest possible AdvertisedWindow allows. 394 5 End-to-End Protocols Bandwidth Time until Wraparound T1 (1.5 Mbps) 6.4 hours Ethernet (10 Mbps) 57 minutes T3 (45 Mbps) 13 minutes FDDI (100 Mbps) 6 minutes STS-3 (155 Mbps) 4 minutes STS-12 (622 Mbps) 55 seconds STS-24 (1.2 Gbps) 28 seconds Table 5.1 Time until 32-bit sequence number space wraps around. Bandwidth Delay × Bandwidth Product T1 (1.5 Mbps) 18 KB Ethernet (10 Mbps) 122 KB T3 (45 Mbps) 549 KB FDDI (100 Mbps) 1.2 MB STS-3 (155 Mbps) 1.8 MB STS-12 (622 Mbps) 7.4 MB STS-24 (1.2 Gbps) 14.8 MB Table 5.2 Required window size for 100-ms RTT. In this case, it is not just the network bandwidth but the delay × bandwidth product that dictates how big the AdvertisedWindow field needs to be—the window needs to be opened far enough to allow a full delay × bandwidth product’s worth of data to be transmitted. Assuming an RTT of 100 ms (a typical number for a cross- country connection in the U.S.), Table 5.2 gives the delay × bandwidth product for several network technologies. As you can see, TCP’s AdvertisedWindow field is in even worse shape than its SequenceNum field—it is not big enough to handle even a T3 connection across the continental United States, since a 16-bit field allows us to advertise a window of only 64 KB. The very same TCP extension mentioned above (see Section 5.2.8) provides a mechanism for effectively increasing the size of the advertised window. 5.2 Reliable Byte Stream (TCP) 395 5.2.5 Triggering Transmission We next consider a surprisingly subtle issue: how TCP decides to transmit a segment. As described earlier,TCP supports a byte-stream abstraction, that is, application programs write bytes into the stream, and it is up to TCP to decide that it has enough bytes to send a segment. What factors govern this decision? If we ignore the possibility of flow control—that is, we assume the window is wide open, as would be the case when a connection first starts—then TCP has three mechanisms to trigger the transmission of a segment. First, TCP maintains a variable, typically called the maximum segment size (MSS), and it sends a segment as soon as it has collected MSS bytes from the sending process. MSS is usually set to the size of the largest segment TCP can send without causing the local IP to fragment. That is, MSS is set to the MTU of the directly connected network, minus the size of the TCP and IP headers. The second thing that triggers TCP to transmit a segment is that the sending process has explicitly asked it to do so. Specifically, TCP supports a push operation, and the sending process invokes this operation to effectively flush the buffer of unsent bytes. The final trigger for transmitting a segment is that a timer fires; the resulting segment contains as many bytes as are currently buffered for transmission. However, as we will soon see, this “timer” isn’t exactly what you expect. Silly Window Syndrome Of course, we can’t just ignore flow control, which plays an obvious role in throttling the sender. If the sender has MSS bytes of data to send and the window is open at least that much, then the sender transmits a full segment. Suppose, however, that the sender is accumulating bytes to send, but the window is currently closed. Now suppose an ACK arrives that effectively opens the window enough for the sender to transmit, say, MSS/2 bytes. Should the sender transmit a half-full segment or wait for the window to open to a full MSS? The original specification was silent on this point, and early implementations of TCP decided to go ahead and transmit a half-full segment. After all, there is no telling how long it will be before the window opens further. It turns out that the strategy of aggressively taking advantage of any available window leads to a situation now known as the silly window syndrome. Figure 5.9 helps visualize what happens. If you think of a TCP stream as a conveyer belt with “full” containers (data segments) going in one direction and empty containers (ACKs) going in the reverse direction, then MSS-sized segments correspond to large containers and 1-byte segments correspond to very small containers. If the sender aggressively fills an empty container as soon as it arrives, then any small container introduced into the system remains in the system indefinitely. That is, it is immediately filled and emptied at each end, and never coalesced with adjacent containers to create larger containers. 396 5 End-to-End Protocols Sender Receiver Figure 5.9 Silly window syndrome. This scenario was discovered when early implementations of TCP regularly found themselves filling the network with tiny segments. Note that the silly window syndrome is only a problem when either the sender transmits a small segment or the receiver opens the window a small amount. If neither of these happens, then the small container is never introduced into the stream. It’s not possible to outlaw sending small segments; for example, the application might do a push after sending a single byte. It is possible, however, to keep the receiver from introducing a small container (i.e., a small open window). The rule is that after advertizing a zero window, the receiver must wait for space equal to an MSS before it advertises an open window. Since we can’t eliminate the possibility of a small container being introduced into the stream, we also need mechanisms to coalesce them. The receiver can do this by delaying ACKs—sending one combined ACK rather than multiple smaller ones—but this is only a partial solution because the receiver has no way of knowing how long it is safe to delay waiting either for another segment to arrive or for the application to read more data (thus opening the window). The ultimate solution falls to the sender, which brings us back to our original issue: When does the TCP sender decide to transmit a segment? Nagle’s Algorithm Returning to the TCP sender, if there is data to send but the window is open less than MSS, then we may want to wait some amount of time before sending the available data, but the question is, how long? If we wait too long, then we hurt interactive applications like Telnet. If we don’t wait long enough, then we risk sending a bunch of tiny packets and falling into the silly window syndrome. The answer is to introduce a timer and to transmit when the timer expires. While we could use a clock-based timer—for example, one that fires every 100 ms—Nagle introduced an elegant self-clocking solution. The idea is that as long as TCP has any data in flight, the sender will eventually receive an ACK. This ACK can be 5.2 Reliable Byte Stream (TCP) 397 treated like a timer firing, triggering the transmission of more data. Nagle’s algorithm provides a simple, unified rule for deciding when to transmit: When the application produces data to send if both the available data and the window ≥ MSS send a full segment else if there is unACKed data in flight buffer the new data until an ACK arrives else send all the new data now In other words, it’s always OK to send a full segment if the window allows. It’s also OK to immediately send a small amount of data if there are currently no segments in transit, but if there is anything in flight, the sender must wait for an ACK before transmiting the next segment. Thus, an interactive application like Telnet that continually writes one byte at a time will send data at a rate of one segment per RTT. Some segments will contain a single byte, while others will contain as many bytes as the user was able to type in one round-trip time. Because some applications cannot afford such a delay for each write they do to a TCP connection, the socket interface allows applications to turn off Nagle’s algorithm by setting the TCP NODELAY option. Setting this option means that data is transmitted as soon as possible. 5.2.6 Adaptive Retransmission Because TCP guarantees the reliable delivery of data, it retransmits each segment if an ACK is not received in a certain period of time. TCP sets this timeout as a function of the RTT it expects between the two ends of the connection. Unfortunately, given the range of possible RTTs between any pair of hosts in the Internet, as well as the varia- tion in RTT between the same two hosts over time, choosing an appropriate timeout value is not that easy. To address this problem, TCP uses an adaptive retransmission mechanism. We now describe this mechanism and how it has evolved over time as the Internet community has gained more experience using TCP. Original Algorithm We begin with a simple algorithm for computing a timeout value between a pair of hosts. This is the algorithm that was originally described in the TCP specification— and the following description presents it in those terms—but it could be used by any end-to-end protocol. The idea is to keep a running average of the RTT and then to compute the timeout as a function of this RTT. Specifically, every time TCP sends a data segment, it records 398 5 End-to-End Protocols the time. When an ACK for that segment arrives, TCP reads the time again and then takes the difference between these two times as a SampleRTT. TCP then computes an EstimatedRTT as a weighted average between the previous estimate and this new sample. That is, EstimatedRTT = α × EstimatedRTT + (1 − α) × SampleRTT The parameter α is selected to smooth the EstimatedRTT. A small α tracks changes in the RTT but is perhaps too heavily influenced by temporary fluctuations. On the other hand, a large α is more stable but perhaps not quick enough to adapt to real changes. The original TCP specification recommended a setting of α between 0.8 and 0.9. TCP then uses EstimatedRTT to compute the timeout in a rather conservative way: TimeOut = 2 × EstimatedRTT Karn/Partridge Algorithm After several years of use on the Internet, a rather obvious flaw was discovered in this simple algorithm. The problem was that an ACK does not really acknowledge a transmission; it actually acknowledges the receipt of data. In other words, whenever a segment is retransmitted and then an ACK arrives at the sender, it is impossible to determine if this ACK should be associated with the first or the second transmission of the segment for the purpose of measuring the sample RTT. It is necessary to know which transmission to associate it with so as to compute an accurate SampleRTT.As illustrated in Figure 5.10, if you assume that the ACK is for the original transmission but it was really for the second, then the SampleRTT is too large (a), while if you assume that the ACK is for the second transmission but it was actually for the first, then the SampleRTT is too small (b). Sender Receiver Original transmission ACK SampleRTT Retransmission Sender Receiver Original transmission ACK SampleRTT Retransmission (a) (b) Figure 5.10 Associating the ACK with (a) original transmission versus (b) retransmission. 5.2 Reliable Byte Stream (TCP) 399 The solution is surprisingly simple. Whenever TCP retransmits a segment, it stops taking samples of the RTT; it only measures SampleRTT for segments that have been sent only once. This solution is known as the Karn/Partridge algorithm, after its inventors. Their proposed fix also includes a second small change to TCP’s timeout mechanism. Each time TCP retransmits, it sets the next timeout to be twice the last timeout, rather than basing it on the last EstimatedRTT. That is, Karn and Partridge proposed that TCP use exponential backoff, similar to what the Ethernet does. The motivation for using exponential backoff is simple: Congestion is the most likely cause of lost segments, meaning that the TCP source should not react too aggressively to a timeout. In fact, the more times the connection times out, the more cautious the source should become. We will see this idea again, embodied in a much more sophisticated mechanism, in Chapter 6. Jacobson/Karels Algorithm The Karn/Partridge algorithm was introduced at a time when the Internet was suffering from high levels of network congestion. Their approach was designed to fix some of the causes of that congestion, and although it was an improvement, the congestion was not eliminated. A couple of years later, two other researchers—Jacobson and Karels— proposed a more drastic change to TCP to battle congestion. The bulk of that proposed change is described in Chapter 6. Here, we focus on the aspect of that proposal that is related to deciding when to time out and retransmit a segment. As an aside, it should be clear how the timeout mechanism is related to congestion—if you time out too soon, you may unnecessarily retransmit a segment, which only adds to the load on the network. As we will see in Chapter 6, the other reason for needing an accurate timeout value is that a timeout is taken to imply conges- tion, which triggers a congestion-control mechanism. Finally, note that there is nothing about the Jacobson/Karels timeout computation that is specific to TCP.It could be used by any end-to-end protocol. The main problem with the original computation is that it does not take the variance of the sample RTTs into account. Intuitively, if the variation among samples is small, then the EstimatedRTT can be better trusted and there is no reason for multi- plying this estimate by 2 to compute the timeout. On the other hand, a large variance in the samples suggests that the timeout value should not be too tightly coupled to the EstimatedRTT. In the new approach, the sender measures a new SampleRTT as before. It then folds this new sample into the timeout calculation as follows: Difference = SampleRTT − EstimatedRTT EstimatedRTT = EstimatedRTT + (δ × Difference) 400 5 End-to-End Protocols Deviation = Deviation + δ(|Difference|−Deviation) where δ is a fraction between 0 and 1. That is, we calculate both the mean RTT and the variation in that mean. TCP then computes the timeout value as a function of both EstimatedRTT and Deviation as follows: TimeOut = μ × EstimatedRTT + φ × Deviation where based on experience, μ is typically set to 1 and φ is set to 4. Thus, when the variance is small, TimeOut is close to EstimatedRTT; a large variance causes the Deviation term to dominate the calculation. Implementation There are two items of note regarding the implementation of timeouts in TCP. The first is that it is possible to implement the calculation for EstimatedRTT and Deviation without using floating-point arithmetic. Instead, the whole calculation is scaled by 2n, with δ selected to be 1/2n. This allows us to do integer arithmetic, implementing multiplication and division using shifts, thereby achieving higher performance. The resulting calculation is given by the following code fragment, where n = 3 (i.e., δ =1/8). Note that EstimatedRTT and Deviation are stored in their scaled-up forms, while the value of SampleRTT at the start of the code and of TimeOut at the end are real, unscaled values. If you find the code hard to follow, you might want to try plugging some real numbers into it and verifying that it gives the same results as the equations above. { SampleRTT -= (EstimatedRTT >> 3); EstimatedRTT += SampleRTT; if (SampleRTT < 0) SampleRTT = -SampleRTT; SampleRTT -= (Deviation >> 3); Deviation += SampleRTT; TimeOut = (EstimatedRTT >> 3) + (Deviation >> 1); } The second point of note is that the Jacobson/Karels algorithm is only as good as the clock used to read the current time. On a typical Unix implementation, the clock granularity is as large as 500 ms, which is significantly larger than the average cross-country RTT of somewhere between 100 and 200 ms. To make matters worse, the Unix implementation of TCP only checks to see if a timeout should happen every time this 500-ms clock ticks, and it only takes a sample of the round-trip time once per 5.2 Reliable Byte Stream (TCP) 401 RTT. The combination of these two factors quite often means that a timeout happens 1 second after the segment was transmitted. Once again, the extensions to TCP include a mechanism that makes this RTT calculation a bit more precise. 5.2.7 Record Boundaries Since TCP is a byte-stream protocol, the number of bytes written by the sender are not necessarily the same as the number of bytes read by the receiver. For example, the application might write 8 bytes, then 2 bytes, then 20 bytes to a TCP connection, while on the receiving side, the application reads 5 bytes at a time inside a loop that iterates 6 times. TCP does not interject record boundaries between the 8th and 9th bytes, nor between the 10th and 11th bytes. This is in contrast to a message-oriented protocol, such as UDP, in which the message that is sent is exactly the same length as the message that is received. Even though TCP is a byte-stream protocol, it has two different features that can be used by the sender to insert record boundaries into this byte stream, thereby informing the receiver how to break the stream of bytes into records. (Being able to mark record boundaries is useful, for example, in many database applications.) Both of these features were originally included in TCP for completely different reasons; they have only come to be used for this purpose over time. The first mechanism is the urgent data feature, as implemented by the URG flag and the UrgPtr field in the TCP header. Originally, the urgent data mechanism was designed to allow the sending application to send out-of-band data to its peer. By “out of band” we mean data that is separate from the normal flow of data (e.g., a command to interrupt an operation already under way). This out-of-band data was identified in the segment using the UrgPtr field and was to be delivered to the receiving process as soon as it arrived, even if that meant delivering it before data with an earlier sequence number. Over time, however, this feature has not been used, so instead of signifying “urgent” data, it has come to be used to signify “special” data, such as a record marker. This use has developed because, as with the push operation, TCP on the receiving side must inform the application that “urgent data” has arrived. That is, the urgent data in itself is not important. It is the fact that the sending process can effectively send a signal to the receiver that is important. The second mechanism for inserting end-of-record markers into a byte is the push operation. Originally, this mechanism was designed to allow the sending process to tell TCP that it should send (flush) whatever bytes it had collected to its peer. The push operation can be used to implement record boundaries because the specification says that TCP must send whatever data it has buffered at the source when the application says push, and optionally, TCP at the destination notifies the application whenever an incoming segment has the PUSH flag set. If the receiving side supports this option 402 5 End-to-End Protocols (the socket interface does not), then the push operation can be used to break the TCP stream into records. Of course, the application program is always free to insert record boundaries without any assistance from TCP. For example, it can send a field that indicates the length of a record that is to follow, or it can insert its own record boundary markers into the data stream. 5.2.8 TCP Extensions We have mentioned at three different points in this section that there are now extensions to TCP that help to mitigate some problem that TCP is facing as the underlying network gets faster. These extensions are designed to have as small an impact on TCP as possible. In particular, they are realized as options that can be added to the TCP header. (We glossed over this point earlier, but the reason that the TCP header has a HdrLen field is that the header can be of variable length; the variable part of the TCP header contains the options that have been added.) The significance of adding these extensions as options rather than changing the core of the TCP header is that hosts can still communicate using TCP even if they do not implement the options. Hosts that do implement the optional extensions, however, can take advantage of them. The two sides agree that they will use the options during TCP’s connection establishment phase. The first extension helps to improve TCP’s timeout mechanism. Instead of mea- suring the RTT using a coarse-grained event, TCP can read the actual system clock when it is about to send a segment, and put this time—think of it as a 32-bit time- stamp—in the segment’s header. The receiver then echoes this timestamp back to the sender in its acknowledgment, and the sender subtracts this timestamp from the cur- rent time to measure the RTT. In essence, the timestamp option provides a convenient place for TCP to “store” the record of when a segment was transmitted; it stores the time in the segment itself. Note that the endpoints in the connection do not need synchronized clocks, since the timestamp is written and read at the same end of the connection. The second extension addresses the problem of TCP’s 32-bit SequenceNum field wrapping around too soon on a high-speed network. Rather than define a new 64-bit sequence number field, TCP uses the 32-bit timestamp just described to effectively extend the sequence number space. In other words, TCP decides whether to accept or reject a segment based on a 64-bit identifier that has the SequenceNum field in the low-order 32 bits and the timestamp in the high-order 32 bits. Since the timestamp is always increasing, it serves to distinguish between two different incarnations of the same sequence number. Note that the timestamp is being used in this setting only to protect against wraparound; it is not treated as part of the sequence number for the purpose of ordering or acknowledging data. 5.2 Reliable Byte Stream (TCP) 403 The third extension allows TCP to advertise a larger window, thereby allow- ing it to fill larger delay × bandwidth pipes that are made possible by high-speed networks. This extension involves an option that defines a scaling factor for the ad- vertised window. That is, rather than interpreting the number that appears in the AdvertisedWindow field as indicating how many bytes the sender is allowed to have unacknowledged, this option allows the two sides of TCP to agree that the Advertised- Window field counts larger chunks (e.g., how many 16-byte units of data the sender can have unacknowledged). In other words, the window scaling option specifies how many bits each side should left-shift the AdvertisedWindow field before using its contents to compute an effective window. 5.2.9 Alternative Design Choices Although TCP has proven to be a robust protocol that satisfies the needs of a wide range of applications, the design space for transport protocols is quite large. TCP is, by no means, the only valid point in that design space. We conclude our discussion of TCP by considering alternative design choices. While we offer an explanation for why TCP’s designers made the choices they did, we leave it to you to decide if there might be a place for alternative transport protocols. First, we have suggested from the very first chapter of this book that there are at least two interesting classes of transport protocols: stream-oriented protocols like TCP and request/reply protocols like RPC. In other words, we have implicitly divided the design space in half and placed TCP squarely in the stream-oriented half of the world. We could further divide the stream-oriented protocols into two groups—reliable and unreliable—with the former containing TCP and the latter being more suitable for interactive video applications that would rather drop a frame than incur the delay associated with a retransmission. This exercise in building a transport protocol taxonomy is interesting and could be continued in greater and greater detail, but the world isn’t as black and white as we might like. Consider the suitability of TCP as a transport protocol for request/reply applications, for example. TCP is a full-duplex protocol, so it would be easy to open a TCP connection between the client and server, send the request message in one direction, and send the reply message in the other direction. There are two com- plications, however. The first is that TCP is a byte-oriented protocol rather than a message-oriented protocol, and request/reply applications always deal with messages. (We explore the issue of bytes versus messages in greater detail in a moment.) The second complication is that in those situations where both the request message and the reply message fit in a single network packet, a well-designed request/reply protocol needs only two packets to implement the exchange, whereas TCP would need at least nine: three to establish the connection, two for the message exchange, and four to tear 404 5 End-to-End Protocols down the connection. Of course, if the request or reply messages are large enough to require multiple network packets (e.g., it might take 100 packets to send a 100,000- byte reply message), then the overhead of setting up and tearing down the connection is inconsequential. In other words, it isn’t always the case that a particular protocol cannot support a certain functionality; it’s sometimes the case that one design is more efficient than another under particular circumstances. Second, as just suggested, you might question why TCP chose to provide a reliable byte-stream service rather than a reliable message-stream service; messages would be the natural choice for a database application that wants to exchange records. There are two answers to this question. The first is that a message-oriented protocol must, by def- inition, establish an upper bound on message sizes. After all, an infinitely long message is a byte stream. For any message size that a protocol selects, there will be applications that want to send larger messages, rendering the transport protocol useless and forcing the application to implement its own transportlike services. The second reason is that while message-oriented protocols are definitely more appropriate for applications that want to send records to each other, you can easily insert record boundaries into a byte stream to implement this functionality, as described in Section 5.2.7. Third, TCP chose to implement explicit setup/teardown phases, but this is not required. In the case of connection setup, it would certainly be possible to send all necessary connection parameters along with the first data message. TCP elected to take a more conservative approach that gives the receiver the opportunity to reject the connection before any data arrives. In the case of teardown, we could quietly close a connection that has been inactive for a long period of time, but this would complicate applications like Telnet that want to keep a connection alive for weeks at a time; such applications would be forced to send out-of-band “keepalive” messages to keep the connection state at the other end from disappearing. Finally, TCP is a window-based protocol, but this is not the only possibility. The alternative is a rate-based design, in which the receiver tells the sender the rate— expressed in either bytes or packets per second—at which it is willing to accept incom- ing data. For example, the receiver might inform the sender that it can accommodate 100 packets a second. There is an interesting duality between windows and rate, since the number of packets (bytes) in the window, divided by the RTT, is exactly the rate. For example, a window size of 10 packets and a 100-ms RTT implies that the sender is allowed to transmit at a rate of 100 packets a second. It is by increasing or decreasing the advertised window size that the receiver is effectively raising or lowering the rate at which the sender can transmit. In TCP, this information is fed back to the sender in the AdvertisedWindow field of the ACK for every segment. One of the key issues in a rate-based protocol is how often the desired rate—which may change over time—is relayed back to the source: Is it for every packet, once per RTT, or only when the 5.3 Remote Procedure Call 405 rate changes? While we have just now considered window versus rate in the context of flow control, it is an even more hotly contested issue in the context of congestion control, which we will discuss in Chapter 6. 5.3 Remote Procedure Call As discussed in Chapter 1, a common pattern of communication used by application programs is the request/reply paradigm, also called message transaction: A client sends a request message to a server, the server responds with a reply message, and the client blocks (suspends execution) waiting for this response. Figure 5.11 illustrates the basic interaction between the client and server in such a message transaction. A transport protocol that supports the request/reply paradigm is much more than a UDP message going in one direction, followed by a UDP message going in the other direction. It also involves overcoming all of the limitations of the underlying network outlined in the problem statement at the beginning of this chapter. While TCP over- comes these limitations by providing a reliable byte-stream service, it doesn’t match the request/reply paradigm very well either since going to the trouble of establishing a TCP connection just to exchange a pair of messages seems like overkill. This section de- scribes a third transport protocol—which we call Remote Procedure Call (RPC)—that more closely matches the needs of an application involved in a request/reply message exchange. RPC is actually more than just a protocol—it is a popular mechanism for struc- turing distributed systems. RPC is popular because it is based on the semantics of a local procedure call—the application program makes a call into a procedure without regard for whether it is local or remote and blocks until the call returns. While this may sound simple, there are two main problems that make RPC more complicated than local procedure calls: Client Server Request Reply Computing Blocked Blocked Blocked Figure 5.11 Timeline for RPC. 406 5 End-to-End Protocols ■ The network between the calling process and the called process has much more complex properties than the backplane of a computer. For example, it is likely to limit message sizes and has a tendency to lose and reorder messages. ■ The computers on which the calling and called processes run may have sig- nificantly different architectures and data representation formats. Thus, a complete RPC mechanism actually involves two major components: 1 A protocol that manages the messages sent between the client and the server pro- cesses and that deals with the potentially undesirable properties of the underlying network 2 Programming language and compiler support to package the arguments into a request message on the client machine and then to translate this message back into the arguments on the server machine, and likewise with the return value (this piece of the RPC mechanism is usually called a stub compiler) Figure 5.12 schematically depicts what happens when a client invokes a remote procedure. First, the client calls a local stub for the procedure, passing it the arguments required by the procedure. This stub hides the fact that the procedure is remote by Caller (client) Client stub RPC protocol Return valueArguments ReplyRequest Callee (server) Server stub RPC protocol Return valueArguments ReplyRequest Figure 5.12 Complete RPC mechanism. 5.3 Remote Procedure Call 407 translating the arguments into a request message and then invoking an RPC protocol to send the request message to the server machine. At the server, the RPC protocol delivers the request message to the server stub, which translates it into the arguments to the procedure and then calls the local procedure. After the server procedure completes, it returns the answer to the server stub, which packages this return value in a reply message that it hands off to the RPC protocol for transmission back to the client. The RPC protocol on the client passes this message up to the client stub, which translates it into a return value that it returns to the client program. This section considers just the protocol-related aspects of an RPC mechanism. That is, it ignores the stubs and focuses instead on the RPC protocol that transmits messages between client and server; the transformation of arguments into messages and vice versa is covered in Chapter 7. Furthermore, since RPC is a generic term— rather than a specific standard like TCP—we are going to take a different approach than we did in the previous section. Instead of organizing the discussion around an existing standard (i.e., TCP) and then pointing out alternative designs at the end, we are going to walk you through the thought process involved in designing an RPC protocol. That is, we will design our own RPC protocol from scratch—considering the design options at every step of the way—and then come back and describe some widely used RPC protocols by comparing and contrasting them to the protocol we just designed. Before jumping in, however, we note that an RPC protocol performs a rather complicated set of functions, and so instead of treating RPC as a single, monolithic protocol, we develop it as a “stack” of three smaller protocols: BLAST, CHAN, and SELECT. Each of these smaller protocols, which we sometimes call a microprotocol, contains a single algorithm that addresses one of the problems outlined at the start of this chapter. As a brief overview: ■ BLAST: fragments and reassembles large messages ■ CHAN: synchronizes request and reply messages ■ SELECT: dispatches request messages to the correct process These microprotocols are complete, self-contained protocols that can be used in dif- ferent combinations to provide different end-to-end services. Section 5.3.4 shows how they can be combined to implement RPC. Just to be clear, BLAST, CHAN, and SELECT are not standard protocols in the sense that TCP, UDP, and IP are. They are simply protocols of our own invention, but ones that demonstrate the algorithms needed to implement RPC. Because this section is not constrained by the artifacts of what has been designed in the past, it provides a particularly good opportunity to examine the principles of protocol design. 408 5 End-to-End Protocols 5.3.1 Bulk Transfer (BLAST) The first problem we are going to tackle is how to turn an underlying network that de- livers messages of some small size (say, 1 KB) into a service that delivers messages of a much larger size (say, 32 KB). While 32 KB does not qualify as “arbitrarily large,” it is large enough to be of practical use for many applications, including most distributed file systems. Ultimately, a stream-based pro- tocol like TCP (see Section 5.2) will be needed to support an arbitrarily large mes- sage, since any message-oriented protocol will necessarily have some upper limit to the size of the message it can handle, and you can always imagine needing to transmit a message that is larger than this limit. We have already examined the basic technique that is used to transmit a large message over a network that can accommo- date only smaller messages—fragmentation and reassembly. We now describe the BLAST protocol, which uses this tech- nique. One of the unique properties of BLAST is how hard it tries to deliver all the fragments of a message. Unlike the AAL segmentation/reassembly mechanism used with ATM (see Section 3.3) or the IP fragmentation/reassembly mechanism (see Section 4.1), BLAST attempts to recover from dropped fragments by retransmitting them. However, BLAST does not go so far as to guarantee message delivery. The sig- nificance of this design choice will become clear later in this section. What Layer Is RPC? Once again, the “What layer is this?” issue raises its ugly head. To many people, especially those who adhere to the Internet archi- tecture, RPC is implemented on top of a transport protocol (usu- ally UDP) and so cannot itself (by definition) be a transport protocol. It is equally valid, however, to ar- gue that the Internet should have an RPC protocol, since it offers a process-to-process service that is fundamentally different from that offered by TCP and UDP. The usual response to such a sugges- tion, however, is that the Internet architecture does not prohibit net- work designers from implementing their own RPC protocol on top of UDP. (In general, UDP is viewed as the Internet architecture’s “es- cape hatch,” since effectively it just adds a layer of demultiplexing to IP.) Whichever side of the issue of whether the Internet should have an official RPC protocol you sup- port, the important point is that the way you implement RPC in the Internet architecture says noth- ing about whether RPC should be BLAST Algorithm The basic idea of BLAST is for the sender to break a large message passed to it by some high-level protocol into a set of smaller fragments, and then for it to transmit 5.3 Remote Procedure Call 409 considered a transport protocol or not. Interestingly, there are other people who believe that RPC is the most interesting protocol in the world and that TCP/IP is just what you do when you want to go “off site.” This is the predominant view of the operating systems commu- nity, which has built countless OS kernels for distributed systems that contain exactly one protocol—you guessed it, RPC—running on top of a network device driver. The water gets even mud- dier when you implement RPC as a combination of three different microprotocols, as is the case in this section. In such a situation, which of the three is the “transport” pro- tocol? Our answer to this ques- tion is that any protocol that offers process-to-process service, as op- posed to node-to-node or host-to- host service, qualifies as a transport protocol. Thus, RPC is a transport protocol and, in fact, can be im- plemented from a combination of microprotocols that are themselves valid transport protocols. these fragments back-to-back over the network. Hence the name BLAST—the pro- tocol does not wait for any of the frag- ments to be acknowledged before send- ing the next. The receiver then sends a selective retransmission request (SRR) back to the sender, indicating which frag- ments arrived and which did not. (The SRR message is sometimes called a par- tial or selective acknowledgment.) Finally, the sender retransmits the missing frag- ments. In the case in which all the frag- ments have arrived, the SRR serves to fully acknowledge the message. Figure 5.13 gives a representative timeline for the BLAST protocol. We now consider the send and re- ceive sides of BLAST in more detail. On the sending side, after fragmenting the mes- sage and transmitting each of the fragments, the sender sets a timer called DONE. When- ever an SRR arrives, the sender retransmits the requested fragments and resets timer DONE. Should the SRR indicate that all the fragments have arrived, the sender frees its copy of the message and cancels timer DONE. If timer DONE ever expires, the sender frees its copy of the message; that is, it gives up. On the receiving side, whenever the first fragment of a message arrives, the re- ceiver initializes a data structure to hold the individual fragments as they arrive and sets a timer LAST FRAG. This timer counts the time that has elapsed since the last fragment arrived. Each time a fragment for that message arrives, the receiver adds it to this data structure, and should all the fragments then be present, it reassembles them into a complete message and passes this message up to the higher-level protocol. There are four exceptional conditions, however, that the receiver watches for: 410 5 End-to-End Protocols Sender Receiver Fragment 1 Fragment 2 Fragment 3 Fragment 5 Fragment 4 Fragment 6 Fragment 3 Fragment 5 SRR SRR Figure 5.13 Representative timeline for BLAST. ■ If the last fragment arrives (the last fragment is specially marked) but the message is not complete, then the receiver determines which fragments are missing and sends an SRR to the sender. It also sets a timer called RETRY. ■ If timer LAST FRAG expires, then the receiver determines which fragments are missing and sends an SRR to the sender. It also sets timer RETRY. ■ If timer RETRY expires for the first or second time, then the receiver deter- mines which fragments are still missing and retransmits an SRR message. ■ If timer RETRY expires for the third time, then the receiver frees the fragments that have arrived and cancels timer LAST FRAG; that is, it gives up. There are three aspects of BLAST worth noting. First, two different events trigger the initial transmission of an SRR: the arrival of the last fragment and the firing of the LAST FRAG timer. In the case of the former, because the network may reorder packets, 5.3 Remote Procedure Call 411 the arrival of the last fragment does not necessarily imply that an earlier fragment is missing (it may just be late in arriving), but since this is the most likely explanation, BLAST aggressively sends an SRR message. In the latter case, we deduce that the last fragment was either lost or seriously delayed. Second, the performance of BLAST does not critically depend on how carefully the timers are set. Timer DONE is used only to decide that it is time to give up and delete the message that is currently being worked on. This timer can be set to a fairly large value since its only purpose is to reclaim storage. Timer RETRY is only used to retransmit an SRR message. Any time the situation is so bad that a protocol is reexecuting a failure recovery process, performance is the last thing on its mind. Finally, timer LAST FRAG has the potential to influence performance—it sometimes triggers the sending by the receiver of an SRR message—but this is an unlikely event: It only happens when the last fragment of the message happens to get dropped in the network. Third, while BLAST is persistent in asking for and retransmitting missing frag- ments, it does not guarantee that the complete message will be delivered. To understand this, suppose that a message consists of only one or two fragments and that these frag- ments are lost. The receiver will never send an SRR, and the sender’s DONE timer will eventually expire, causing the sender to release the message. To guarantee deliv- ery, BLAST would need for the sender to time out if it does not receive an SRR and then retransmit the last set of fragments it had transmitted. While BLAST certainly could have been designed to do this, we chose not to because the purpose of BLAST is to deliver large messages, not to guarantee message delivery. Other protocols can be configured on top of BLAST to guarantee message delivery. You might wonder why we put any retransmission capability at all into BLAST if we need to put a guaran- teed delivery mechanism above it anyway. The reason is that we’d prefer to retransmit only those fragments that were lost rather than having to retransmit the entire larger message whenever one fragment is lost. So we get the guarantees from the higher-level protocol but some improved efficiency by retransmitting fragments in BLAST. BLAST Message Format The BLAST header has to convey several pieces of information. First, it must contain some sort of message identifier so that all the fragments that belong to the same message can be identified. Second, there must be a way to identify where in the original message the individual fragments fit, and likewise, an SRR must be able to indicate which fragments have arrived and which are missing. Third, there must be a way to distinguish the last fragment, so that the receiver knows when it is time to check to see if all the fragments have arrived. Finally, it must be possible to distinguish a data 412 5 End-to-End Protocols Data ProtNum MID Length NumFrags Type FragMask 01631 Figure 5.14 Format for BLAST message header. message from an SRR message. Some of these items are encoded in a header field in an obvious way, but others can be done in a variety of different ways. Figure 5.14 gives the header format used by BLAST. The following discussion explains the various fields and considers alternative designs. The MID field uniquely identifies this message. All fragments that belong to the same message have the same value in their MID field. The only question is how many bits are needed for this field. This is similar to the question of how many bits are needed in the SequenceNum field for TCP. The central issue in deciding how many bits to use in the MID field has to do with how long it will take before this field wraps around and the protocol starts using message ids over again. If this happens too soon—that is, the MID field is only a few bits long—then it is possible for the protocol to become confused by a message that was delayed in the network, so that an old incarnation of some message id is mistaken for a new incarnation of that same id. So, how many bits are enough to ensure that the amount of time it takes for the MID field to wrap around is longer than the amount of time a message can potentially be delayed in the network? In the worst-case scenario, each BLAST message contains a single fragment that is 1 byte long, which means that BLAST might need to generate a new MID for every byte it sends. On a 10-Mbps Ethernet, this would mean generating a new MID roughly once every microsecond, while on a 1.2-Gbps STS-24 link, a new MID would be required once every 7 nanoseconds. Of course, this is a ridiculously conservative calculation— the overhead involved in preparing a message is going to be more than a microsecond. Thus, suppose a new MID is potentially needed once every microsecond, and a mes- sage may be delayed in the network for up to 60 seconds (our standard worst-case 5.3 Remote Procedure Call 413 assumption for the Internet); then we need to ensure that there are more than 60 million MID values. While a 26-bit field would be sufficient (226 = 67,108,864), it is easier to deal with header fields that are even multiples of a byte, so we will settle on a 32-bit MID field.◮ This conservative (you could say paranoid) analysis of the MID field illustrates an important point. When designing a transport protocol, it is tempting to take shortcuts, since not all networks suffer from all the problems listed in the problem statement at the beginning of this chapter. For example, messages do not get stuck in an Ethernet for 60 seconds, and similarly, it is physically impossible to reorder messages on an Ethernet segment. The problem with this way of thinking, however, is that if you want the transport protocol to work over any kind of network, then you have to design for the worst case. This is because the real danger is that as soon as you assume that an Ethernet does not reorder packets, someone will come along and put a bridge or a router in the middle of it. Let’s move on to the other fields in the BLAST header. The Type field indicates whether this is a DATA message or an SRR message. Notice that while we certainly don’t need 16 bits to represent these two types, as a general rule we like to keep the header fields aligned on 32-bit (word) boundaries, so as to improve processing efficiency. The ProtNum field identifies the high-level protocol that is configured on top of BLAST; incoming messages are demultiplexed to this protocol. The Length field indicates how many bytes of data are in this fragment; it has nothing to do with the length of the entire message. The NumFrags field indicates how many fragments are in this message. This field is used to determine when the last fragment has been received. An alternative is to include a flag that is only set for the last fragment. Finally, the FragMask field is used to distinguish among fragments. It is a 32-bit field that is used as a bit mask. For messages of Type = DATA, the ith bit is 1 (all others are 0) to indicate that this message carries the ith fragment. For messages of Type = SRR, the ith bit is set to 1 to indicate that the ith fragment has arrived, and it is set to 0 to indicate that the ith fragment is missing. Note that there are several ways to identify fragments. For example, the header could have contained a simple “frag- ment ID” field, with this field set to i to denote the ith fragment. The tricky part with this approach, as opposed to a bit-vector, is how the SRR specifies which fragments have arrived and which have not. If it takes an n-bit number to identify each missing fragment—as opposed to a single bit in a fixed-size bit-vector—then the SRR message will be of variable length, depending on how many fragments are missing. Variable- length headers are allowed, but they are a little trickier to process. On the other hand, one limitation of the BLAST header given above is that the length of the bit-vector limits each message to only 32 fragments. If the underlying network has an MTU of 1 KB, then this is sufficient to send up to 32-KB messages. 414 5 End-to-End Protocols 5.3.2 Request/Reply (CHAN) The next microprotocol, CHAN, implements the request/reply algorithm that is at the core of RPC. In terms of the common properties of transport protocols given in the problem statement at the beginning of this chapter, CHAN guarantees message de- livery, ensures that only one copy of each message is delivered, and allows the com- municating processes to synchronize with each other. In the case of this last item, the synchronization we are after mimics the behavior of a procedure call—the caller (client) blocks while waiting for a reply from the callee (server). At-Most-Once Semantics The name CHAN comes from the fact that the protocol implements a logical request/ reply channel between a pair of participants. At any given time, there can be only one message transaction active on a given channel. Like the concurrent logical channel protocol described in Section 2.5.3, the application programs have to open multiple channels if they want to have more than one request/reply transaction between them at the same time. The most important property of each channel is that it preserves a semantics known as at-most-once. This means that for every request message that the client sends, at most one copy of that message is delivered to the server. Stated in terms of the RPC mechanism that CHAN is designed to support, for each time the client calls a remote procedure, that procedure is invoked at most one time on the server machine. We say “at-most-once” rather than “exactly once” because it is always possible that either the network or the server machine has failed, making it impossible to deliver even one copy of the request message. As obvious as at-most-once sounds, not all RPC protocols support this behavior. Some support a semantics that is facetiously called zero-or-more semantics, that is, each invocation on a client results in the remote procedure being invoked zero or more times. It is not hard to understand how this would cause problems for a remote procedure that changed some local state variable (e.g., incremented a counter) or that had some externally visible side effect (e.g., launched a missile) each time it was invoked. On the other hand, if the remote procedure being invoked is idempotent— multiple invocations have the same effect as just one—then the RPC mechanism need not support at-most-once semantics; a simpler (possibly faster) implementation will suffice. CHAN Algorithm The request/reply algorithm has several subtle aspects; hence, we develop it in stages. The basic algorithm is straightforward, as illustrated by the timeline given in Figure 5.15. The client sends a request message and the server acknowledges it. Then, after 5.3 Remote Procedure Call 415 Client Server Request ACK Reply ACK Figure 5.15 Simple timeline for CHAN. Client Server Request 1 Request 2 Reply 2 Reply 1 … Figure 5.16 Timeline for CHAN when using implicit ACKs. executing the procedure, the server sends a reply message and the client acknowledges the reply. Because the reply message often comes back with very little delay, and it is some- times the case that the client turns around and makes a second request on the same channel immediately after receiving the first reply, this basic scenario can be optimized by using a technique called implicit acknowledgments. As illustrated in Figure 5.16, the reply message serves to acknowledge the request message, and a subsequent request acknowledges the preceding reply. 416 5 End-to-End Protocols There are two factors that potentially complicate the rosy picture we have painted so far. The first is that either a message carrying data (a request message or a reply message) or the ACK sent to acknowledge that message may be lost in the network. To account for this possibility, both client and server save a copy of each message they send until an ACK for it has arrived. Each side also sets a RETRANSMIT timer and resends the message should this timer expire. Both sides reset this timer and try again some agreed-upon number of times before giving up and freeing the message. Recall from Section 2.5.1 that this acknowledgment/timeout strategy means that it is possible for duplicate copies of a message to arrive—the original message arrives, the ACK is lost, and then the retransmission arrives. Thus, the receiver must remember what messages it has seen and discard any duplicates. This is done through the use of a MID field in the header. Any message whose MID field does not match the next expected MID is discarded instead of being passed up to the high-level protocol configured on top of CHAN. The second complication is that the server may take an arbitrarily long time to produce the result, and worse yet, it may crash (either the process or the entire machine) before generating the reply. Keep in mind that we are talking about the period of time after the server has acknowledged the request but before it has sent the reply. To help the client distinguish between a slow server and a dead server, CHAN’s client side periodically sends an “Are you alive?” message to the server, and CHAN’s server side responds with an ACK. Alternatively, the server could send “I am still alive” messages to the client without the client having first solicited them, but we prefer the client-initiated approach because it keeps the server as simple as possible (i.e., it has one less timer to manage). CHAN Message Format The CHAN message format is given in Figure 5.17. As with BLAST, the Type field specifies the type of the message; in this case, the possible types are REQ, REP, ACK, and PROBE.(PROBE is the “Are you alive?” message discussed above.) Similarly, the ProtNum field identifies the high-level protocol that depends on CHAN. The CID field uniquely identifies the logical channel to which this message belongs. This is a 16-bit field, meaning that CHAN supports up to 64K concurrent request/reply transactions between any pair of hosts. Of course, a given host can be participating in channels with many other hosts at the same time. The MID field uniquely identifies each request/reply pair; the reply message has the same MID as the request. Note that because CHAN permits only one message transaction at a time on a given channel, you might think that a 1-bit MID field is sufficient, just as for the stop-and-wait algorithm presented in Section 2.5.1. However, 5.3 Remote Procedure Call 417 Data Type MID BID Length ProtNum CID 01631 Figure 5.17 Format for CHAN message header. as with BLAST, we have to be concerned about messages that wander around the network for an extended period of time and then suddenly appear at the destination, confusing CHAN. Thus, using much the same reasoning as we used in Section 5.3.1, CHAN uses a 32-bit MID field. Finally, the BID field gives the boot id for the host. A machine’s boot id is a number that is incremented each time the machine reboots; this number is read from disk, incremented, and written back to disk during the machine’s startup procedure. This number is then put in every message sent by that host. The role played by the BID field is much the same as the role played by the large MID field—it protects against old messages suddenly appearing at the destination—although in this case, the old message is due not to an arbitrary delay in the network but rather to a machine that has crashed and rebooted. To understand the use of the boot id, consider the following pathological sit- uation. A client machine sends a request message with MID = 0, then crashes and reboots, and then sends an unrelated request message, also with MID = 0. The server may not have been aware that the client crashed and rebooted, and upon seeing a re- quest message with MID = 0, acknowledges it and discards it as a duplicate. To protect against this possibility, each side of CHAN makes sure that the BID, MID pair, not just the MID, matches what it is expecting. BID is also a 32-bit field, which means that if we assume that it takes at least 10 minutes to reboot a machine, it will wrap around once every 40 billion minutes (approximately 80,000 years). In effect, the BID and MID combine to form a unique 64-bit id for each transaction; the low-order 32 bits are incremented for each transaction but reset to 0 when the machine reboots, and the high-order 32 bits are incremented each time the machine reboots. 418 5 End-to-End Protocols Timeouts CHAN involves three different timers: There is a RETRANSMIT timer on both the client and server, and the client also manages a PROBE timer. The PROBE timer is not critical to performance and thus can be set to a conservatively large value—on the order of several seconds. The RETRANSMIT timer, however, does influence the performance of CHAN. If it is set too large, then CHAN might wait an unnecessarily long time before retransmitting a message that was lost by the network. This clearly hurts performance. If the RETRANSMIT timer is set too small, however, then CHAN may load the network with unnecessary traffic. If CHAN is designed to run on a local area network only, or even over a campus- size extended LAN, then RETRANSMIT can be set to a fixed value. Something on the order of 20 milliseconds would be reasonable. This is because the RTT of a LAN is not that variable. If CHAN is expected to run over the Internet, however, then selecting a suitable value for RETRANSMIT is similar to the problem faced by TCP. Thus, CHAN would calculate the RETRANSMIT timeout using a mechanism similar to the one described in Section 5.2.6. The only difference is that CHAN has to take into account the fact that the message it is sending ranges in size from 1 byte to 32 KB, whereas TCP is always transmitting segments of approximately the same size. Synchronous versus Asynchronous Protocols One way to characterize a protocol is by whether it is synchronous or asynchronous. These two terms have significantly different meanings, depending on where in the protocol hierarchy you use them. At the transport layer, it is most accurate to think of synchrony as a spectrum of possibilities rather than as two alternatives, where the key attribute of any point along the spectrum is how much the sender knows, after the operation to send a message returns. In other words, if we assume that an application program invokes a send operation on a transport protocol, then the question is, Exactly what does the application know about the success of the operation when the send operation returns? At the asynchronous end of the spectrum, the application knows absolutely noth- ing when send returns. It not only doesn’t know if the message was received by its peer, but it doesn’t even know for sure that the message has successfully left the local ma- chine. At the synchronous end of the spectrum, the send operation typically returns a reply message. That is, the application not only knows that the message it sent was re- ceived by its peer, but it knows that the peer has returned an answer. Thus, synchronous protocols implement the request/reply abstraction, while asynchronous protocols are used if the sender wants to be able to transmit many messages without having to wait for a response. Using this definition, CHAN is obviously a synchronous protocol. 5.3 Remote Procedure Call 419 Although we have not discussed them in this chapter, there are interesting points between these two extremes. For example, the transport protocol might implement send so that it blocks (does not return) until the message has been successfully received at the remote machine, but returns before the sender’s peer on that machine has actually processed and responded to it. This is sometimes called a reliable datagram protocol. Implementation of CHAN We conclude our discussion of CHAN by giving fragments of C code that implement its client side. Reading code can be tedious, but if done judiciously, it can help to solidify your understanding of how a system works. In the case of CHAN, it serves to illustrate all the separate pieces that go into a protocol implementation—the function that sends an outgoing message, the function that retransmits messages, and the function that processes incoming messages—and how they interact with each other. We begin with CHAN’s two key data structures: ChanHdr and ChanState. The fields in ChanHdr have already been explained. The fields in ChanState will be explained by the code that follows. Note that ChanState includes a hdr template field, which is a copy of the CHAN header. Many of the fields in the CHAN header remain the same for all messages sent out over this channel. These fields are filled in when the channel is created (not shown); only the fields that change are modified before a given message is transmitted. typedef struct { u_short Type; /* message type: REQ, REP, ACK, PROBE */ u_short CID; /* unique channel id */ int MID; /* unique message id */ int BID; /* unique boot id */ int Length; /* length of message */ int ProtNum; /* high-level protocol number */ } ChanHdr; typedef struct { u_char type; /* type of session: CLIENT or SERVER */ u_char status; /* status of session: BUSY or IDLE */ Event event; /* place to save timeout event */ int timeout; /* timeout value */ int retries; /* number of times retransmitted */ int ret_val; /* place to save return value */ Msg *request; /* place to save request message */ Msg *reply; /* place to save reply message */ Semaphore reply_sem; /* semaphore the client blocks on */ int mid; /* message id for this channel */ int bid; /* boot id for this channel */ 420 5 End-to-End Protocols ChanHdr hdr_template; /* header template for this channel */ BlastState blast; /* pointer to BLAST protocol */ } ChanState; We now turn our attention to the function that sends request messages. Since CHAN exports a synchronous interface to higher-level protocols—the caller blocks un- til a reply can be returned—the send operation we have been assuming since Chapter 1 is not going to work. Therefore, we introduce a new interface operation, which we give the generic name call, that blocks until a reply message is available, and re- turns that reply message to the caller. The first argument identifies the channel being used; it effectively encapsulates all the information needed to send the message to the correct destination. The second and third arguments correspond to the abstract data type (ADT) for messages, and represent the request and reply messages, respec- tively. We assume this ADT supports the obvious operations (e.g., msgSaveCopy and msgLength). int callCHAN(ChanState *state, Msg *request, Msg *reply) { ChanHdr *hdr; char hbuf[HLEN]; /* ensure only one transaction per channel */ if ((state->status != IDLE)) return FAILURE; state->status = BUSY; /* save a copy of request msg and pointer to reply msg*/ msgSaveCopy(&state->request, request); state->reply = reply; /* fill out header fields */ hdr = state->hdr_template; hdr->Length = msgLength(request); if (state->mid == MAX_MID) state->mid = 0; hdr->MID = ++state->mid; /* attach header to msg and send it */ store_chan_hdr(hdr, hbuf); msgAddHdr(request, hdr, HLEN); sendBLAST(request); /* schedule first timeout event */ state->retries = 1; state->event = eventSchedule(retransmit, state, state->timeout); 5.3 Remote Procedure Call 421 /* block waiting for the reply msg */ semWait(&state->reply_sem); /* clean up state and return */ flush_msg(state->request); state->status = IDLE; return state->ret_val; } The first thing to notice is that the ChanState passed as an argument to callCHAN includes a field named status that indicates whether or not this channel is being used. If the channel is currently in use, then callCHAN returns failure. An alternative design would be to block the calling thread until the channel becomes idle. We have elected to push responsibility for blocking threads that want to use busy channels onto the higher-level protocol, in our case, SELECT. The next thing to notice about call is that after filling out the message header and transmitting the request message via BLAST, the calling process is blocked on a semaphore (reply sem). When the reply message eventually arrives, it is processed by CHAN’s deliverCHAN routine (see below), which copies the reply message into state variable reply and signals this blocked process. The process then returns. Should the reply message not arrive, then timeout routine retransmit is called (see below). This event is scheduled in the body of callCHAN. The next routine, retransmit, is called whenever the retransmit timer fires. It is scheduled for the first time in callCHAN, but each time it is called, it reschedules itself. Once the request message has been retransmitted four times, CHAN gives up: It sets the return value to FAILURE and wakes up the blocked client process. Finally, each time re- transmit executes and sends another copy of the request message, it needs to resave the message in state variable request. This is because we assume that each time a protocol sends a message to the lower-level protocol, it loses its reference to the message. static void retransmit(Event ev, int *arg) { ChanState *state = (ChanState *)arg; Msg tmp; /* unblock the client process if we have retried 4 times */ if (++state->retries > 4) { state->ret_val = FAILURE; semSignal(state->rep_sem); return; } /* retransmit request message */ 422 5 End-to-End Protocols msgSaveCopy(&tmp, &state->request); sendBLAST(&tmp); /* reschedule event with exponential backoff */ state->timeout = 2*state->timeout; state->event = eventSchedule(retransmit, state, state->timeout); } Finally, we consider CHAN’s deliver routine. The first thing we observe is that CHAN is an asymmetric protocol: The code that implements CHAN on the client machine is completely distinct from the code that implements CHAN on the server machine. This fact is stored in the CHAN state variable (type). Thus, the first thing CHAN’s deliver routine does is check to see whether it is running on a server (i.e., it expects REQ messages) or on a client (i.e., it expects REP messages), and then it calls the appropriate client- or server-specific routine. In this case, we show the client-specific routine, deliverClient. static int deliverClient(ChanState *state, Msg *msg) { ChanHdr hdr; char *hbuf; /* strip header and verify correctness */ hbuf = msgStripHdr(msg, HLEN); load_chan_hdr(&hdr, hbuf); if (!clnt_msg_ok(state, &hdr)) return FAILURE; /* cancel retransmit timeout event */ eventCancel(state->event); /* if this is an ACK, then schedule PROBE timer and exit*/ if (hdr.Type == ACK) { state->event = eventSchedule(probe, s, PROBE); return SUCCESS; } /* msg is a REP; save it and signal blocked client */ msgSaveCopy(state->reply, msg); state->ret_val = SUCCESS; semSignal(&state->reply_sem); return SUCCESS; } 5.3 Remote Procedure Call 423 Routine deliverClient first checks to see if it has received the expected message, for example, that it has the right MID, the right BID, and that the message is of type REP or ACK. This check is made in subroutine clnt msg ok (not shown). If it is a valid acknowledgment message, then deliverClient cancels the RETRANSMIT timer and schedules the PROBE timer. The PROBE timer is not shown, but would be similar to the RETRANSMIT timer given above. If the message is a valid reply, then deliverClient cancels the RETRANSMIT timer, saves a copy of the reply message in state variable reply, and wakes up the blocked client process. It is this client process that actually returns the reply message to the high-level protocol; the process that called deliverClient simply returns back down the protocol stack. 5.3.3 Dispatcher (SELECT) The final microprotocol, called SELECT, dispatches request messages to the appro- priate procedure. It is the RPC protocol stack’s version of a demultiplexing protocol like UDP; the main difference is that it is a synchronous protocol rather than an asyn- chronous protocol. What this means is that on the client side, SELECT is given a proce- dure number that the client wants to invoke, it puts this number in its header, and then it invokes the call operation on a lower-level request/reply protocol like CHAN. When this invocation returns, SELECT merely lets the return pass through to the client; it has no real demultiplexing work to do. On the server side, SELECT uses the procedure number it finds in its header to select the right local procedure to invoke. When this pro- cedure returns, SELECT simply returns to the low-level protocol that just invoked it. It may seem that SELECT is so simple that it is not worthy of being treated as a separate protocol. After all, CHAN already has its own demultiplexing field that could be used to dispatch incoming request messages to the appropriate procedure. There are two reasons why we elected to separate SELECT into a self-contained protocol. The first is that doing so makes it possible to change the address space with which procedures are identified simply by configuring a different version of SELECT into the protocol graph. In some settings, it is sufficient to define a flat address space for procedures—for example, a 16-bit selector field allows you to identify 64K different procedures. In other settings, however, a flat address space is hard to manage—who decides which procedure gets which procedure number? In this case, it might be better to have a hierarchical address space, that is, a two-part procedure number. First, each program could be given a program number, where a program corresponds to something like a “file server” or a “name server.” Next, each program could be given the responsibility to assign unique procedure numbers to its own procedures. For example, within the file server program, read might be procedure 1, write might be procedure 2, seek might be procedure 3, and so on, whereas within the name server program, insert might be procedure 1 and lookup might be procedure 2. 424 5 End-to-End Protocols The second reason we implement SELECT as its own protocol is that it provides a good place to manage concurrency. Recall that CHAN supports at-most-once chan- nels. Suppose we want to allow applications running on this host to make multiple outstanding calls to the same remote procedure. Since CHAN allows only one out- standing call at a time, the only way to do this is to open multiple channels to the same server. Each time a calling process invokes SELECT, it sends the process out on an idle channel. If all the channels are currently active, then SELECT blocks the calling process until a channel becomes idle. 5.3.4 Putting It All Together (SunRPC, DCE) We are now ready to construct an RPC stack from the microprotocols described in the three previous subsections. This section also explains two widely used RPC protocols— SunRPC and DCE-RPC—in terms of our three microprotocols. A Simple RPC Stack Figure 5.18 depicts a simple protocol stack that implements RPC. At the bottom are the protocols that implement the underlying network. Although this stack could contain protocols corresponding to any of the networking technologies discussed in the three previous chapters, we use IP running on top of an Ethernet for illustrative purposes. On top of IP is BLAST, which turns the small message size of the underlying network into a communication service that supports messages of up to 32 KB in length. Notice that it is not strictly true that the underlying network provides for only small messages; IP can handle messages of up to 64 KB. However, because IP BLAST ETH IP SELECT CHAN Figure 5.18 A simple RPC stack. 5.3 Remote Procedure Call 425 has to fragment such large messages before sending them out over the Ethernet, and BLAST’s fragmentation/reassembly algorithm is superior to IP’s (because it is able to selectively retransmit missing fragments), we prefer to treat IP as though it supports ex- actly the same MTU as the underlying physical network. This puts the fragmentation/ reassembly burden on BLAST, unless IP has to perform fragmentation out in the middle of the network somewhere. Next, CHAN implements the request/reply algorithm. Recall that we chose not to implement reliable delivery in BLAST, but instead postponed solving this issue until a higher-level protocol. In this case, CHAN’s timeout and acknowledgment mecha- nism makes sure messages are reliably delivered. Other protocols might use different techniques to guarantee delivery or, for that matter, might choose not to implement reliable delivery at all. This is an example of the end-to-end argument at work—do not do at low levels of the system (e.g., BLAST) what has to be done at higher levels (e.g., CHAN) anyway. Finally, SELECT defines an address space for identifying remote procedures. As suggested in Section 5.3.3, different versions of SELECT, each defining a different method for identifying procedures, could be configured on top of CHAN. In fact, it would even be possible to write a version of SELECT that mimics some existing RPC package’s address space for procedures (such as SunRPC’s), and then to use CHAN and BLAST underneath this new SELECT to implement the rest of the RPC stack. This new stack would not interoperate with the original protocol, but it would allow you to slide a new RPC system underneath an existing collection of remote procedures without having to change the interface. SELECT also manages concurrency. SunRPC SELECT, CHAN, and BLAST, although complete and correctly functioning protocols, have been neither standardized nor widely adopted. We now turn our discussion to a widely used RPC protocol—SunRPC. Ironically, SunRPC has also not been approved by any standardization body, but it has become a de facto standard, thanks to its wide distribution with Sun workstations and to the central role it plays in Sun’s popular Network File System (NFS). At the time of this writing, the IETF is considering officially adopting SunRPC as a standard Internet protocol. Fundamentally, any RPC protocol must worry about three issues: fragment- ing large messages, synchronizing request and reply messages, and dispatching re- quest messages to the appropriate procedure. SunRPC is no exception. Unlike the SELECT/CHAN/BLAST stack, however, SunRPC addresses these three functions in a different order and uses slightly different algorithms. The basic SunRPC protocol graph is given in Figure 5.19. 426 5 End-to-End Protocols IP ETH SunRPC UDP Figure 5.19 Protocol graph for SunRPC. First, SunRPC implements the core request/reply algorithm; it is CHAN’s coun- terpart. SunRPC differs from CHAN, however, in that it does not technically guarantee at-most-once semantics; there are obscure circumstances under which a duplicate copy of a request message is delivered to the server (see below). Second, the role of SELECT is split between UDP and SunRPC—UDP dispatches to the correct program, and SunRPC dispatches to the correct procedure within the program. (We discuss how procedures are identified in more detail below.) Finally, the ability to send request and reply messages that are larger than the network MTU, corresponding to the func- tionality implemented in BLAST, is handled by IP. Keep in mind, however, that IP is not as persistent as BLAST is in implementing fragmentation; BLAST uses selective retransmission, whereas IP does not. As just mentioned, SunRPC uses two-tier addresses to identify remote proce- dures: a 32-bit program number and a 32-bit procedure number. (There is also a 32-bit version number, but we ignore that in the following discussion.) For example, the NFS server has been assigned program number x00100003, and within this pro- gram, getattr is procedure 1, setattr is procedure 2, read is procedure 6, write is procedure 8, and so on. Each program is reachable by sending a message to some UDP port. When a request message arrives at this port, SunRPC picks it up and calls the appropriate procedure. To determine which port corresponds to a particular SunRPC program number, there is a separate SunRPC program, called the Port Mapper, that maps program num- bers to port numbers. The Port Mapper itself also has a program number (x00100000) that must be translated into some UDP port. Fortunately, the Port Mapper is always present at a well-known UDP port (111). The Port Mapper program supports several procedures, one of which (procedure number 3) is the one that performs the program- to-port number mapping. 5.3 Remote Procedure Call 427 Thus, to send a request message to NFS’s read procedure, a client first sends a request message to the Port Mapper at well-known UDP port 111, asking that pro- cedure 3 be invoked to map program number x00100003 to the UDP port where the NFS program currently resides. (In practice, NFS is such an important program that it is given its own well-known UDP port, so the Port Mapper need not be involved in finding it.) The client then sends a SunRPC request message with procedure number 6 to this UDP port, and the SunRPC module listening at that port calls the NFS read procedure. The client also caches the program-to-port number mapping so that it need not go back to the Port Mapper each time it wants to talk to the NFS program. The actual SunRPC header is defined by a complex nesting of data structures. Figure 5.20 gives the essential details for the case in which the call completes with- out any problems. XID is a unique transaction id, much like CHAN’s MID field. The reason that SunRPC cannot guarantee at-most-once semantics is that on the server side, SunRPC does not remember that it has already seen a particular XID once it has successfully completed the transaction. This is only a problem if the client retransmits a request message as a result of a timeout and that request message is in transit at exactly the same time as the reply to the original request is on its way from the server back to the client. When the retransmitted request arrives at the server, it looks like a new transaction, since the server thinks it has already completed the transaction with Data Data MsgType = CALL XID RPCVersion = 2(a) Program Version Procedure Credentials (variable) Verifier (variable) MsgType = REPLY XID Status = ACCEPTED(b) 031031 Figure 5.20 SunRPC header formats: (a) request; (b) reply. 428 5 End-to-End Protocols this XID. Clearly, if the reply arrives at the client before the timeout, then the request will not be retransmitted. Likewise, if the retransmitted request arrives at the server before the reply has been generated, then the server will recognize that transaction XID is already in progress, and it will discard the duplicate request message. So it is really quite unlikely that this erroneous behavior will occur. Note that the server’s short-term memory about XIDs also means that it cannot protect itself against messages that have been delayed for a long time in the network. This has not been a serious problem with SunRPC, however, because it was originally designed for use on a LAN. Returning to the SunRPC header format, the request message contains variable- length Credentials and Verifier fields, both of which are used by the client to authenticate itself to the server, that is, to give evidence that the client has the right to invoke the server. How a client authenticates itself to a server is a general issue that must be addressed by any protocol that wants to provide a reasonable level of security. This topic is discussed in more detail in Chapter 8. DCE The Distributed Computing Environment (DCE) defines another widely used RPC protocol, which we call DCE-RPC. DCE is a set of standards and software for build- ing distributed systems. It was defined by the Open Software Foundation (OSF), a consortium of computer companies that originally included IBM, Digital, and Hewlett-Packard; today OSF goes by the name Open Group. DCE-RPC is the RPC protocol at the core of the DCE system. It can be used with the Network Data Representation (NDR) stub compiler described in Chapter 7, but it also serves as the underlying RPC protocol for the Common Object Request Broker Architecture (CORBA), which is an industrywide standard for building distributed, object-oriented systems. DCE-RPC is designed to run on top of UDP. It is similar to SunRPC in that it defines a two-level addressing scheme: UDP demultiplexes to the correct server, DCE- RPC dispatches to a particular procedure exported by that server, and clients consult an “endpoint mapping service” (similar to SunRPC’s Port Mapper) to learn how to reach a particular server. Unlike SunRPC, however, DCE-RPC implements at-most-once call semantics. It does this in a single protocol that essentially combines the algorithms in BLAST and CHAN. We focus our discussion on this aspect of DCE-RPC. (In truth, DCE-RPC supports multiple call semantics, including an idempotent semantics similar to SunRPC’s, but at-most-once is the default behavior.) Figure 5.21 gives a timeline for the typical exchange of messages, where each message is labelled by its DCE-RPC type. The pattern is similar to CHAN’s: The client sends a Request message, the server eventually replies with a Response message, and 5.3 Remote Procedure Call 429 Client Server Request Ack Response Ping Working Ping Working Figure 5.21 Typical DCE-RPC message exchange. the client acknowledges (Ack) the response. Instead of the server acknowledging the request messages, however, the client periodically sends a Ping message to the server, which responds with a Working message to indicate that the remote procedure is still in progress. Although not shown in the figure, other message types are also supported. For example, the client can send a Quit message to the server,asking it to abort an earlier call that is still in progress; the server responds with a Quack (quit acknowledgment) message. Also, the server can respond to a Request message with a Reject message (indicating that a call has been rejected), and it can respond to a Ping message with a Nocall message (indicating that the server has never heard of the caller). In addition to the message type, request and reply messages include four key fields that are used to implement both the fragmentation/reassembly aspects of BLAST and the message transaction aspects of CHAN. These include ServerBoot, ActivityId, SequenceNum, and FragmentNum. The ServerBoot field serves the same purpose as CHAN’s BID (boot id) field: The server records its boot time in a global variable each time it starts up, and it includes this 430 5 End-to-End Protocols variable in each call it services. The ActivityId field is similar to CHAN’s CID (channel id) field: It identifies a logical connection between the client and server on which a sequence of calls can be made. The SequenceNum field then distinguishes between calls made as part of the same activity; it serves the same purpose as CHAN’s MID (message id) and SunRPC’s xid (transaction id) fields. Like CHAN (and unlike SunRPC), DCE-RPC keeps track of the last sequence number used as part of a particular activity, so as to ensure at-most-once semantics. Because both request and response messages may be larger than the underlying network packet size, they may be fragmented into multiple packets. The FragmentNum field uniquely identifies each fragment that makes up a given request or reply message. Unlike BLAST, which uses a bit-vector to identify fragments, each DCE-RPC fragment is assigned a unique fragment number (e.g., 0, 1, 2, 3, and so on). Both the client and server implement a selective acknowledgment mechanism, which works as follows. (We describe the mechanism in terms of a client sending a fragmented request message to the server; the same mechanism applies when a server sends a fragment response to the client.) First, each fragment that makes up the request message contains both a unique FragmentNum and a flag indicating whether this packet is a fragment of a call (frag)or the last fragment of a call (last frag); request messages that fit in a single packet carry a no frag flag. The server knows it has received the complete request message when it has the last frag packet and there are no gaps in the fragment numbers. Second, in response to each arriving fragment, the server sends a Fack (fragment acknowledgment) message to the client. This acknowledgment identifies the highest fragment number that the server has successfully received. In other words, the acknowledgment is cumulative, much like in TCP. In addition, however, the server selectively acknowledges any higher fragment numbers it has received out of order. It does so with a bit-vector that identifies these out-of-order fragments relative to the highest in-order fragment it has received. Finally, the client responds by retransmitting the missing fragments. Figure 5.22 illustrates how this all works. Suppose the server has successfully received fragments up through number 20, plus fragments 23, 25, and 26. The server responds with a Fack that identifies fragment 20 as the highest in-order fragment, plus a bit-vector (SelAck) with the third (23 = 20 + 3), fifth (25 = 20 + 5), and sixth (26 = 20 + 6) bits turned on. So as to support an (almost) arbitrarily long bit-vector, the size of the vector (measured in 32-bit words) is given in the SelAckLen field. Given DCE-RPC’s support for very large messages—the FragmentNum field is 16 bits long, meaning it can support 64K fragments—it is not appropriate for the protocol to blast all the fragments that make up a message as fast as it can, 5.4 Performance 431 Client Server Fack 19 20 21 22 23 24 25 26 21 22 24 Type = Fack FragmentNum = 20 WindowSize = 10 SelAckLen = 1 SelAck[1] = 0x36 110100 6 + 20 5 + 20 3 + 20 FragmentNum … … … … Figure 5.22 Fragmentation with selective acknowledgments. as BLAST does, since doing so might overrun the receiver. Instead, DCE-RPC im- plements a flow-control algorithm that is very similar to TCP’s. Specifically, each Fack message not only acknowledges received fragments, but also informs the sender of how many fragments it may now send. This is the purpose of the WindowSize field in Figure 5.22, which serves exactly the same purpose as TCP’s Advertised- Window field except it counts fragments rather than bytes. DCE-RPC also imple- ments a congestion-control mechanism that is similar to TCP’s, which we will see in Chapter 6. 5.4 Performance Recall that Chapter 1 introduced the two quantitative metrics by which network performance is evaluated: latency and throughput. As mentioned in that discussion, these metrics are influenced not only by the underlying hardware (e.g., propagation delay and link bandwidth) but also by software overheads. Now that we have a complete software-based protocol graph available to us that includes alternative transport protocols, we can discuss how to meaningfully measure its performance. 432 5 End-to-End Protocols Linux kernel Linux kernel 100 Mbps User process User process Figure 5.23 Measured system: Two Pentium workstations running Linux connected by a 100-Mbps Ethernet. The importance of such measurements is that they represent the performance seen by application programs. We begin, as any report of experimental results should, by describing our ex- perimental method. This includes the apparatus used in the experiments. We ran our experiments on a pair of 733-MHz Pentium workstations connected by an isolated 100-Mbps Ethernet. The Ethernet spanned a single machine room so propagation is not an issue, making this a measure of processor/software overheads. Each worksta- tion was running the Linux operating system (2.4 kernel). A test program running on top of the socket interface ping-pongs (reflects) messages between the two machines. Figure 5.23 illustrates one round-trip between the two test programs. Each experiment involved running three identical instances of the same test. Each test, in turn, involved sending a message of some specified size back and forth between the two machines 10,000 times. The system’s clock was read at the beginning and end of each test, and the difference between these two times was divided by 10,000 to determine the time taken for each round-trip. The average of these three times (the three runs of the test) is reported for each experiment below. Each experiment involved a different-sized message. The latency numbers were for message sizes of 1 byte, 100 bytes, 200 bytes, ..., 1000 bytes. The throughput results were for message sizes of 1 KB, 2 KB, 4 KB, 8 KB, ..., 32 KB. Table 5.3 gives the results of the latency test. As you would expect, latency increases with message size. Although there are sometimes special cases where you might be interested in the latency of, say, a 200-byte message, typically the most 5.4 Performance 433 Message Size (Bytes) UDP TCP 1 58 66 100 76 84 200 93 104 300 111 124 400 132 136 500 150 159 600 167 176 700 184 194 800 203 210 900 223 228 1000 239 249 Table 5.3 Measured round-trip latencies (µs) for various message sizes and protocols. important latency number is the 1-byte case. This is because the 1-byte case represents the overhead involved in processing each message that does not depend on the amount of data contained in the message. It is typically the lower bound on latency, representing factors like the speed-of-light delay and the time taken to process headers. Note that there is a small difference between the latency experienced by the two different protocol stacks, with UDP round-trip times a bit less than for TCP. This is to be expected since TCP provides more functionality. The results of the throughput test are given in Figure 5.24. Here, we show only the results for UDP. The key thing to notice in this graph is that throughput improves as the messages get larger. This makes sense—each message involves a certain amount of overhead, so a larger message means that this overhead is amortized over more bytes. The throughput curve flattens off above 16 KB, at which point the per-message overhead becomes insignificant when compared to the large number of bytes that the protocol stack has to process. A second thing to notice is that the throughput curve tops out before reaching 100 Mbps. Although it can’t be deduced from these measurements, it turns out that the factor preventing our system from running at the full Ethernet speed is a limitation of the network adaptor rather than the software. 434 5 End-to-End Protocols 1 2 4 8 16 32 64 Message size (KB) Measured throughput (Mbps) 100 70 75 80 85 90 95 Figure 5.24 Measured throughput using UDP, for various message sizes. 5.5 Summary This chapter has described three very different end-to-end protocols. The first protocol we considered is a simple demultiplexer, as typified by UDP. All such a protocol does is dispatch messages to the appropriate application process based on a port number. It does not enhance the best-effort service model of the underlying network in any way, or said another way, it offers an unreliable, connectionless datagram service to application programs. The second type is a reliable byte-stream protocol, and the specific example of this type that we looked at is TCP. The challenges faced with such a protocol are to recover from messages that may be lost by the network, to deliver messages in the same order in which they are sent, and to allow the receiver to do flow control on the sender. TCP uses the basic sliding window algorithm, enhanced with an advertised window, to implement this functionality. The other item of note for this protocol is the importance of an accurate timeout/retransmission mechanism. Interestingly, even though TCP is a single protocol, we saw that it employs at least five different algorithms—sliding Open Issue: Application-Specific Protocols 435 window, Nagle, three-way handshake, Karn/Partridge, and Jacobson/Karels—all of which can be of value to any end-to-end protocol. The third transport protocol we looked at is a request/reply protocol that forms the basis for RPC. In this case, a combination of three different algorithms are em- ployed to implement the request/reply service: a selective retransmission algorithm that is used to fragment and reassemble large messages, a synchronous channel algorithm that pairs the request message with the reply message, and a dispatch algorithm that causes the correct remote procedure to be invoked. What should be clear after reading this chapter is that transport proto- col design is a tricky business. As we have seen, getting a transport pro- tocol right in the first place is hard enough, but changing circumstances make matters more complicated. The OPEN ISSUE Application-Specific Protocols challenge is finding ways to adapt to these changes. Our experience with using the protocol can change. As we saw with TCP’s time- out mechanism, experience led to a series of refinements in how TCP decides to retrans- mit a segment. None of these changes affected the format of the TCP header, however, and so they could be incorporated into TCP one implementation at a time. That is, there was no need for everyone to upgrade their version of TCP on the same day. The characteristics of the underlying network can also change. For many years, TCP’s 32-bit sequence number and 16-bit advertised window were more than ade- quate. Recently, however, higher-bandwidth networks have meant that the sequence number is not large enough to protect against wraparound, and the advertised window is too small to allow the sender to fill the network pipe. While an obvious solution would have been to redefine the TCP header to include a 64-bit sequence number field and a 32-bit advertised window field, this would have introduced the very serious prob- lem of how several million Internet hosts would make the transition from the current header to this new header. While such transitions have been performed on production networks, including the telephone network, they are no trivial matter. It was decided, therefore, to implement the necessary extensions as options and to allow hosts to nego- tiate with each other as to whether or not they will use the options for each connection. This approach will not work indefinitely, however,since the TCP header has room for only 44 bytes of options. (This is because the HdrLen field is 4 bits long, meaning that the total TCP header length cannot exceed 16× 32 bit words, or 64 bytes.) Of course, a TCP option that extends the space available for options is always a possibility, but you have to wonder how far it is worth going for the sake of backward compatibility. 436 5 End-to-End Protocols Perhaps the hardest changes to accommodate are the adaptations to the level of service required by application programs. It is inevitable that some applications will have a good reason for wanting a slight variation from the standard services. For ex- ample, some applications want RPC most of the time, but occasionally want to be able to send a stream of request messages without waiting for any of the replies. While this is no longer technically the semantics of RPC, a common scenario is to modify an existing RPC protocol to allow this flexibility. As another example, because video is a stream-oriented application, it is tempting to use TCP as the transport protocol. Unfortunately, TCP guarantees reliability, which is not important to the video applica- tion. In fact, a video application would rather drop a frame (segment) than wait for it to be retransmitted. Rather than invent a new transport protocol from scratch, how- ever, some designers have proposed that TCP should support an option that effectively turns off its reliability feature. It seems that such a protocol could hardly be called TCP anymore, but we are talking about the pragmatics of getting an application to run. How to develop transport protocols that can evolve to satisfy diverse applica- tions, many of which have not yet been imagined, is a hard problem. It is possible that the ultimate answer to this problem is the one-function-per-protocol style exem- plified by the microprotocols we used to implement RPC, or some similar mechanism by which the application programmer is allowed to program, configure, or otherwise customize the transport protocol. FURTHER READING There is no doubt that TCP is a complex protocol and that in fact it has subtleties not illuminated in this chapter. Therefore, the recommended reading list for this chapter includes the original TCP specification. Our motivation for including this specification is not so much to fill in the missing details as to expose you to what an honest-to- goodness protocol specification looks like. The other two papers in the recommended reading list focus on RPC. The paper by Birrell and Nelson is the seminal paper on the topic, while the article by O’Malley and Peterson describes the one-function-per- protocol design philosophy in more detail. ■ USC-ISI. Transmission Control Protocol. Request for Comments 793, Septem- ber 1981. ■ Birrell, A., and B. Nelson. Implementing remote procedure calls. ACM Trans- actions on Computer Systems 2(1):39–59, February 1984. ■ O’Malley, S., and L. Peterson. A dynamic network architecture. ACM Trans- actions on Computer Systems 10(2):110–143, May 1992. Exercises 437 Beyond the protocol specification, the most complete description of TCP, includ- ing its implementation in Unix, can be found in [Ste94b] and [SW95]. Also, the third volume of Comer and Stevens’s TCP/IP series of books describes how to write client/server applications on top of TCP and UDP, using the Posix socket interface [CS00], the Windows socket interface [CS97], and the System V Unix TLI interface [CS94]. Several papers evaluate the performance of different transport protocols at a very detailed level. For example, the article by Clark et al. [CJRS89] measures the processing overheads of TCP, a paper by Mosberger et al. [MPBO96] explores the limitations of protocol processing overheads, and Thekkath and Levy [TL93] and Schroeder and Burrows [SB89] examine RPC’s performance in great detail. The original TCP timeout calculation was described in the TCP specification (see above), while the Karn/Partridge algorithm was described in [KP91] and the Jacobson/ Karels algorithm was proposed in [Jac88]. The TCP extensions are defined by Jacobson et al. [JBB92], while O’Malley and Peterson [OP91] argue that extending TCP in this way is not the right approach to solving the problem. Finally, there are several distributed operating systems that have defined their own RPC protocol. Notable examples include the V system, described by Cheriton and Zwaenepoel [CZ85]; Sprite, described by Ousterhout et al. [OCD+88]; and Amoeba, described by Mullender [Mul90]. The latest version of SunRPC, as defined by Srini- vasan [Sri95a], is a proposed standard for the Internet. EXERCISES 1 If a UDP datagram is sent from host A, port P to host B, port Q, but at host B there is no process listening to port Q, then B is to send back an ICMP Port Unreachable message to A. Like all ICMP messages, this is addressed to A as a whole, not to portPonA. (a) Give an example of when an application might want to receive such ICMP messages. (b) Find out what an application has to do, on the operating system of your choice, to receive such messages. (c) Why might it not be a good idea to send such messages directly back to the originating port P on A? 2 Consider a simple UDP-based protocol for requesting files (based somewhat loosely on the Trivial File Transport Protocol, TFTP). The client sends an 438 5 End-to-End Protocols initial file request, and the server answers (if the file can be sent) with the first data packet. Client and server then continue with a stop-and-wait transmission mechanism. (a) Describe a scenario by which a client might request one file but get another; you may allow the client application to exit abruptly and be restarted with the same port. (b) Propose a change in the protocol that will make this situation much less likely. 3 Design a simple UDP-based protocol for retrieving files from a server. No authen- tication is to be provided. Stop-and-wait transmission of the data may be used. Your protocol should address the following issues: (a) Duplication of the first packet should not duplicate the “connection.” (b) Loss of the final ACK should not necessarily leave the server in doubt as to whether the transfer succeeded. (c) A late-arriving packet from a past connection shouldn’t be interpretable as part of a current connection. 4 This chapter explains three sequences of state transitions during TCP connec- tion teardown. There is a fourth possible sequence, which traverses an additional arc (not shown in Figure 5.7) from FIN WAIT 1 to TIME WAIT and labelled FIN + ACK/ACK. Explain the circumstances that result in this fourth teardown sequence. 5 When closing a TCP connection, why is the two-segment-lifetime timeout not necessary on the transition from LAST ACK to CLOSED? 6 A sender on a TCP connection that receives a 0 advertised window periodically probes the receiver to discover when the window becomes nonzero. Why would the receiver need an extra timer if it were responsible for reporting that its advertised window had become nonzero (i.e., if the sender did not probe)? 7 Read the man page (or Windows equivalent) for the Unix/Windows utility netstat. Use netstat to see the state of the local TCP connections. Find out how long closing connections spend in TIME WAIT. 8 The sequence number field in the TCP header is 32 bits long, which is big enough to cover over 4 billion bytes of data. Even if this many bytes were never transferred Exercises 439 over a single connection, why might the sequence number still wrap around from 232 − 1to0? 9 You are hired to design a reliable byte-stream protocol that uses a sliding window (like TCP). This protocol will run over a 100-Mbps network. The RTT of the network is 100 ms, and the maximum segment lifetime is 60 seconds. (a) How many bits would you include in the AdvertisedWindow and Sequence- Num fields of your protocol header? (b) How would you determine the numbers given above, and which values might be less certain? 10 You are hired to design a reliable byte-stream protocol that uses a sliding window (like TCP). This protocol will run over a 1-Gbps network. The RTT of the network is 140 ms, and the maximum segment lifetime is 60 seconds. How many bits would you include in the AdvertisedWindow and SequenceNum fields of your protocol header? 11 Suppose a host wants to establish the reliability of a link by sending packets and measuring the percentage that are received; routers, for example, do this. Explain the difficulty of doing this over a TCP connection. 12 Suppose TCP operates over a 1-Gbps link. (a) Assuming TCP could utilize the full bandwidth continuously, how long would it take the sequence numbers to wrap around completely? (b) Suppose an added 32-bit timestamp field increments 1000 times during the wraparound time you found above. How long would it take for the timestamp to wrap around? 13 Suppose TCP operates over a 40-Gbps STS-768 link. (a) Assuming TCP could utilize the full bandwidth continuously, how long would it take the sequence numbers to wrap around completely? (b) Suppose an added 32-bit timestamp field increments 1000 times during the wraparound time you found above. How long would it take for the timestamp to wrap around? 14 If host A receives two SYN packets from the same port from remote host B, the second may be either a retransmission of the original or else, if B has crashed and rebooted, an entirely new connection request. 440 5 End-to-End Protocols (a) Describe the difference as seen by host A between these two cases. (b) Give an algorithmic description of what the TCP layer needs to do upon receiving a SYN packet. Consider the duplicate/new cases above, and the pos- sibility that nothing is listening to the destination port. 15 Suppose x and y are two TCP sequence numbers. Write a function to determine whether x comes before y (in the notation of Request for Comments 793, “x =< y”) or after y; your solution should work even when sequence numbers wrap around. 16 Suppose an idle TCP connection exists between sockets A and B. A third party has eavesdropped and knows the current sequence number at both ends. (a) Suppose the third party sends A a forged packet ostensibly from B and with 100 bytes of new data. What happens? Hint: Look up in Request for Com- ments 793 what TCP does when it receives an ACK that is not an “accept- able ACK.” (b) Suppose the third party sends each end such a forged 100-byte data packet ostensibly from the other end. What happens now? What would happen if A later sent 200 bytes of data to B? 17 Suppose party A connects to the Internet via a dial-up IP server (e.g., using SLIP or PPP), has several open Telnet connections (using TCP), and is cut off. Party B then dials in and is assigned the same IP address that A had. Assuming B was able to guess to what host(s) A had been connected, describe a sequence of probes that could enable B to obtain sufficient state information to continue with A’s connections. 18 Diagnostic programs are commonly available that record the first 100 bytes, say, of every TCP connection to a certain host, port. Outline what must be done with each received TCP packet, P, in order to determine if it contains data that belongs to the first 100 bytes of a connection to host HOST, port PORT. Assume the IP header is P.IPHEAD, the TCP header is P.TCPHEAD, and header fields are as named in Figures 4.3 and 5.4. Hint: To get initial sequence numbers (ISNs) you will have to examine every packet with the SYN bit set. Ignore the fact that sequence numbers will eventually be reused. 19 If a packet arrives at host A with B’s source address, it could just as easily have been forged by any third host C. If, however, A accepts a TCP connection from B, then during the three-way handshake A sent ISNA to B’s address and received an Exercises 441 acknowledgment of it. If C is not located so as to be able to eavesdrop on ISNA, then it might seem that C could not have forged B’s response. However, the algorithm for choosing ISNA does give other unrelated hosts a fair chance of guessing it. Specifically, A selects ISNA based on a clock value at the time of connection. Request for Comments 793 specifies that this clock value be incremented every 4 μs; common Berkeley implementations once simplified this to incrementing by 250,000 (or 256,000) once per second. (a) Given this simplified increment-once-per-second implementation, explain how an arbitrary host C could masquerade as B in at least the opening of a TCP connection. You may assume that B does not respond to SYN + ACK packets A is tricked into sending to it. (b) Assuming real RTTs can be estimated to within 40 ms, about how many tries would you expect it to take to implement the strategy of part (a) with the unsimplified “increment every 4 μs” TCP implementation? 20 The Nagle algorithm, built into most TCP implementations, requires the sender to hold a partial segment’s worth of data (even if PUSHed) until either a full segment accumulates or the most recent outstanding ACK arrives. (a) Suppose the letters abcdefghi are sent, one per second, over a TCP connection with an RTT of 4.1 seconds. Draw a timeline indicating when each packet is sent and what it contains. (b) If the above were typed over a full-duplex Telnet connection, what would the user see? (c) Suppose that mouse position changes are being sent over the connection. As- suming that multiple position changes are sent each RTT, how would a user perceive the mouse motion with and without the Nagle algorithm? 21 Suppose a client C repeatedly connects via TCP to a given port on a server S, and that each time it is C that initiates the close. (a) How many TCP connections a second can C make here before it ties up all its available ports in TIME WAIT state? Assume client ephemeral ports are in the range 1024–5119, and that TIME WAIT lasts 60 seconds. (b) Berkeley-derived TCP implementations typically allow a socket in TIME WAIT state to be reopened before TIME WAIT expires, if the highest se- quence number used by the old incarnation of the connection is less than the ISN used by the new incarnation. This solves the problem of old data accepted as new; however, TIME WAIT also serves the purpose of handling late final 442 5 End-to-End Protocols FINs. What would such an implementation have to do to address this and still achieve strict compliance with the TCP requirement that a FIN sent anytime before or during a connection’s TIME WAIT receive the same response? 22 Explain why TIME WAIT is a somewhat more serious problem if the server ini- tiates the close than if the client does. Describe a situation in which this might reasonably happen. 23 What is the justification for the exponential increase in timeout value proposed by Karn and Partridge? Why, specifically, might a linear (or slower) increase be less desirable? 24 The Jacobson/Karels algorithm sets TimeOut to be 4 mean deviations above the mean. Assume that individual packet round-trip times follow a statistical normal distribution, for which 4 mean deviations are π standard deviations. Using statis- tical tables, for example, what is the probability that a packet will take more than TimeOut time to arrive? 25 Suppose a TCP connection, with window size 1, loses every other packet. Those that do arrive have RTT = 1 second. What happens? What happens to TimeOut? Do this for two cases: (a) After a packet is eventually received, we pick up where we left off, resuming with EstimatedRTT initialized to its pretimeout value and TimeOut double that. (b) After a packet is eventually received, we resume with TimeOut initialized to the last exponentially backed-off value used for the timeout interval. In the following four exercises, the calculations involved are straightforward with a spreadsheet. 26 Suppose, in TCP’s adaptive retransmission mechanism, that EstimatedRTT is 4.0 at some point and subsequent measured RTTs all are 1.0. How long does it take before the TimeOut value, as calculated by the Jacobson/Karels algorithm, falls below 4.0? Assume a plausible initial value of Deviation; how sensitive is your answer to this choice? Use δ = 1/8. 27 Suppose, in TCP’s adaptive retransmission mechanism, that EstimatedRTT is 90 at some point and subsequent measured RTTs all are 200. How long does it take before the TimeOut value, as calculated by the Jacobson/Karels algorithm, falls below 300? Assume initial Deviation value of 25; use δ = 1/8. Exercises 443 28 Suppose TCP’s measured RTT is 1.0 except that every Nth RTT is 4.0. What is the largest N, approximately, that doesn’t result in timeouts in the steady state (i.e., for which the Jacobson/Karels TimeOut remains greater than 4.0)? Use δ = 1/8. 29 Suppose that TCP is measuring RTTs of 1.0 second, with a mean deviation of 0.1 second. Suddenly the RTT jumps to 5.0 seconds, with no deviation. Compare the behaviors of the original and Jacobson/Karels algorithms for computing Time- Out. Specifically, how many timeouts are encountered with each algorithm? What is the largest TimeOut calculated? Use δ = 1/8. 30 Suppose that, when a TCP segment is sent more than once, we take SampleRTT to be the time between the original transmission and the ACK, as in Figure 5.10(a). Show that if a connection with a 1-packet window loses every other packet (i.e., each packet is transmitted twice), then EstimatedRTT increases to infinity. As- sume Tim