path: root/doc/rfc/rfc3819.txt

Network Working Group                                       P. Karn, Ed.
Request for Comments: 3819                                      Qualcomm
BCP: 89                                                       C. Bormann
Category: Best Current Practice                  Universitaet Bremen TZI
                                                            G. Fairhurst
                                                  University of Aberdeen
                                                             D. Grossman
                                                          Motorola, Inc.
                                                               R. Ludwig
                                                       Ericsson Research
                                                              J. Mahdavi
                                                                  Novell
                                                           G. Montenegro
                                   Sun Microsystems Laboratories, Europe
                                                                J. Touch
                                                                 USC/ISI
                                                                 L. Wood
                                                           Cisco Systems
                                                               July 2004


                Advice for Internet Subnetwork Designers

Status of this Memo

   This document specifies an Internet Best Current Practices for the
   Internet Community, and requests discussion and suggestions for
   improvements.  Distribution of this memo is unlimited.

Copyright Notice

   Copyright (C) The Internet Society (2004).

Abstract

   This document provides advice to the designers of digital
   communication equipment, link-layer protocols, and packet-switched
   local networks (collectively referred to as subnetworks), who wish to
   support the Internet protocols but may be unfamiliar with the
   Internet architecture and the implications of their design choices on
   the performance and efficiency of the Internet.










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Table of Contents

   1.  Introduction and Overview. . . . . . . . . . . . . . . . . . .  2
   2.  Maximum Transmission Units (MTUs) and IP Fragmentation . . . .  4
       2.1.  Choosing the MTU in Slow Networks. . . . . . . . . . . .  6
   3.  Framing on Connection-Oriented Subnetworks . . . . . . . . . .  7
   4.  Connection-Oriented Subnetworks. . . . . . . . . . . . . . . .  9
   5.  Broadcasting and Discovery . . . . . . . . . . . . . . . . . . 10
   6.  Multicasting . . . . . . . . . . . . . . . . . . . . . . . . . 11
   7.  Bandwidth on Demand (BoD) Subnets. . . . . . . . . . . . . . . 13
   8.  Reliability and Error Control. . . . . . . . . . . . . . . . . 14
       8.1.  TCP vs Link-Layer Retransmission . . . . . . . . . . . . 14
       8.2.  Recovery from Subnetwork Outages . . . . . . . . . . . . 17
       8.3.  CRCs, Checksums and Error Detection. . . . . . . . . . . 18
       8.4.  How TCP Works. . . . . . . . . . . . . . . . . . . . . . 20
       8.5.  TCP Performance Characteristics. . . . . . . . . . . . . 22
             8.5.1.  The Formulae . . . . . . . . . . . . . . . . . . 22
             8.5.2.  Assumptions. . . . . . . . . . . . . . . . . . . 23
             8.5.3.  Analysis of Link-Layer Effects on TCP
                     Performance. . . . . . . . . . . . . . . . . . . 24
   9.  Quality-of-Service (QoS) Considerations. . . . . . . . . . . . 26
   10. Fairness vs Performance. . . . . . . . . . . . . . . . . . . . 29
   11. Delay Characteristics. . . . . . . . . . . . . . . . . . . . . 30
   12. Bandwidth Asymmetries. . . . . . . . . . . . . . . . . . . . . 31
   13. Buffering, Flow and Congestion Control . . . . . . . . . . . . 31
   14. Compression. . . . . . . . . . . . . . . . . . . . . . . . . . 34
   15. Packet Reordering. . . . . . . . . . . . . . . . . . . . . . . 36
   16. Mobility . . . . . . . . . . . . . . . . . . . . . . . . . . . 37
   17. Routing. . . . . . . . . . . . . . . . . . . . . . . . . . . . 39
   18. Security Considerations. . . . . . . . . . . . . . . . . . . . 41
   19. Contributors . . . . . . . . . . . . . . . . . . . . . . . . . 44
   20. Informative References . . . . . . . . . . . . . . . . . . . . 45
   21. Contributors' Addresses. . . . . . . . . . . . . . . . . . . . 57
   22. Authors' Addresses . . . . . . . . . . . . . . . . . . . . . . 58
   23. Full Copyright Statement . . . . . . . . . . . . . . . . . . . 60

1.  Introduction and Overview

   IP, the Internet Protocol [RFC791] [RFC2460], is the core protocol of
   the Internet.  IP defines a simple "connectionless" packet-switched
   network.  The success of the Internet is largely attributed to IP's
   simplicity, the "end-to-end principle" [SRC81] on which the Internet
   is based, and the resulting ease of carrying IP on a wide variety of
   subnetworks, not necessarily designed with IP in mind.  A subnetwork
   refers to any network operating immediately below the IP layer to
   connect two or more systems using IP (i.e., end hosts or routers).
   In its simplest form, this may be a direct connection between the IP
   systems (e.g., using a length of cable or a wireless medium).



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   This document defines a subnetwork as a layer 2 network, which is a
   network that does not rely upon the services of IP routers to forward
   packets between parts of the subnetwork.  However, IP routers may
   bridge frames at Layer 2 between parts of a subnetwork.  Sometimes,
   it is convenient to aggregate a group of such subnetworks into a
   single logical subnetwork.  IP routing protocols (e.g., OSPF, IS-IS,
   and PIM) can be configured to support this aggregation, but typically
   present a layer-3 subnetwork rather than a layer-2 subnetwork.  This
   may also result in a specific packet passing several times over the
   same layer-2 subnetwork via an intermediate layer-3 gateway (router).
   Because that aggregation requires layer-3 components, issues thereof
   are beyond the scope of this document.

   However, while many subnetworks carry IP, they do not necessarily do
   so with maximum efficiency, minimum complexity, or cost, nor do they
   implement certain features to efficiently support newer Internet
   features of increasing importance, such as multicasting or quality of
   service.

   With the explosive growth of the Internet, IP packets comprise an
   increasingly large fraction of the traffic carried by the world's
   telecommunications networks.  It therefore makes sense to optimize
   both existing and new subnetwork technologies for IP as much as
   possible.

   Optimizing a subnetwork for IP involves three complementary
   considerations:

   1.  Providing functionality sufficient to carry IP.

   2.  Eliminating unnecessary functions that increase cost or
       complexity.

   3.  Choosing subnetwork parameters that maximize the performance of
       the Internet protocols.

   Because IP is so simple, consideration 2 is more of an issue than
   consideration 1.  That is to say, subnetwork designers make many more
   errors of commission than errors of omission.  However, certain
   enhancements to Internet features, such as multicasting and quality-
   of-service, benefit significantly from support given by the
   underlying subnetworks beyond that necessary to carry "traditional"
   unicast, best-effort IP.








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   A major consideration in the efficient design of any layered
   communication network is the appropriate layer(s) in which to
   implement a given function.  This issue was first addressed in the
   seminal paper, "End-to-End Arguments in System Design" [SRC81].  That
   paper argued that many functions can be implemented properly *only*
   on an end-to-end basis, i.e., at the highest protocol layers, outside
   the subnetwork.  These functions include ensuring the reliable
   delivery of data and the use of cryptography to provide
   confidentiality and message integrity.

   Such functions cannot be provided solely by the concatenation of
   hop-by-hop services; duplicating these functions at the lower
   protocol layers (i.e., within the subnetwork) can be needlessly
   redundant or even harmful to cost and performance.

   However, partial duplication of functionality in a lower layer can
   *sometimes* be justified by performance, security, or availability
   considerations.  Examples include link-layer retransmission to
   improve the performance of an unusually lossy channel, e.g., mobile
   radio, link-level encryption intended to thwart traffic analysis, and
   redundant transmission links to improve availability, increase
   throughput, or to guarantee performance for certain classes of
   traffic.  Duplication of protocol functions should be done only with
   an understanding of system-level implications, including possible
   interactions with higher-layer mechanisms.

   The original architecture of the Internet was influenced by the
   end-to-end principle [SRC81], and has been, in our view, part of the
   reason for the Internet's success.

   The remainder of this document discusses the various subnetwork
   design issues that the authors consider relevant to efficient IP
   support.

2.  Maximum Transmission Units (MTUs) and IP Fragmentation

   IPv4 packets (datagrams) vary in size, from 20 bytes (the size of the
   IPv4 header alone) to a maximum of 65535 bytes.  Subnetworks need not
   support maximum-sized (64KB) IP packets, as IP provides a scheme that
   breaks packets that are too large for a given subnetwork into
   fragments that travel as independent IP packets and are reassembled
   at the destination.  The maximum packet size supported by a
   subnetwork is known as its Maximum Transmission Unit (MTU).

   Subnetworks may, but are not required to, indicate the length of each
   packet they carry.  One example is Ethernet with the widely used DIX
   [DIX82] (not IEEE 802.3 [IEEE8023]) header, which lacks a length




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   field to indicate the true data length when the packet is padded to a
   minimum of 60 bytes.  This is not a problem for uncompressed IP
   because each IP packet carries its own length field.

   If optional header compression [RFC1144] [RFC2507] [RFC2508]
   [RFC3095] is used, however, it is required that the link framing
   indicate frame length because that is needed for the reconstruction
   of the original header.

   In IP version 4 (the version now in widespread use), fragmentation
   can occur at either the sending host or in an intermediate router,
   and fragments can be further fragmented at subsequent routers if
   necessary.

   In IP version 6 [RFC2460], fragmentation can occur only at the
   sending host; it cannot occur in a router (called "router
   fragmentation" in this document).

   Both IPv4 and IPv6 provide a "path MTU discovery" procedure [RFC1191]
   [RFC1435] [RFC1981] that allows the sending host to avoid
   fragmentation by discovering the minimum MTU along a given path and
   reduce its packet sizes accordingly.  This procedure is optional in
   IPv4 and IPv6.

   Path MTU discovery is widely deployed, but it sometimes encounters
   problems.  Some routers fail to generate the ICMP messages that
   convey path MTU information to the sender, and sometimes the ICMP
   messages are blocked by overly restrictive firewalls.  The result can
   be a "Path MTU Black Hole" [RFC2923] [RFC1435].

   The Path MTU Discovery procedure, the persistence of path MTU black
   holes, and the deletion of router fragmentation in IPv6 reflect a
   consensus of the Internet technical community that router
   fragmentation is best avoided.  This requires that subnetworks
   support MTUs that are "reasonably" large.  All IPv4 end hosts are
   required to accept and reassemble IP packets of size 576 bytes
   [RFC791], but such a small value would clearly be inefficient.
   Because IPv6 omits fragmentation by routers, [RFC2460] specifies a
   larger minimum MTU of 1280 bytes.  Any subnetwork with an internal
   packet payload smaller than 1280 bytes must implement a mechanism
   that performs fragmentation/reassembly of IP packets to/from
   subnetwork frames if it is to support IPv6.

   If a subnetwork cannot directly support a "reasonable" MTU with
   native framing mechanisms, it should internally fragment.  That is,
   it should transparently break IP packets into internal data elements
   and reassemble them at the other end of the subnetwork.




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   This leaves the question of what is a "reasonable" MTU.  Ethernet (10
   and 100 Mb/s) has an MTU of 1500 bytes, and because of the ubiquity
   of Ethernet few Internet paths currently have MTUs larger than this
   value.  This severely limits the utility of larger MTUs provided by
   other subnetworks.  Meanwhile, larger MTUs are increasingly desirable
   on high-speed subnetworks to reduce the per-packet processing
   overhead in host computers, and implementers are encouraged to
   provide them even though they may not be usable when Ethernet is also
   in the path.

   Various "tunneling" schemes, such as GRE [RFC2784] or IP Security in
   tunnel mode [RFC2406], treat IP as a subnetwork for IP.  Since
   tunneling adds header overhead, it can trigger fragmentation, even
   when the same physical subnetworks (e.g., Ethernet) are used on both
   sides of the host performing IPsec encapsulation.  Tunneling has made
   it more difficult to avoid router fragmentation and has increased the
   incidence of path MTU black holes [RFC2401] [RFC2923].  Larger
   subnetwork MTUs may help to alleviate this problem.

2.1.  Choosing the MTU in Slow Networks

   In slow networks, the largest possible packet may take a considerable
   amount of time to send.  This is known as channelisation or
   serialisation delay.  Total end-to-end interactive response time
   should not exceed the well-known human factors limit of 100 to 200
   ms.  This includes all sources of delay: electromagnetic propagation
   delay, queuing delay, serialisation delay, and the store-and-forward
   time, i.e., the time to transmit a packet at link speed.

   At low link speeds, store-and-forward delays can dominate total
   end-to-end delay; these are in turn directly influenced by the
   maximum transmission unit (MTU) size.  Even when an interactive
   packet is given a higher queuing priority, it may have to wait for a
   large bulk transfer packet to finish transmission.  This worst-case
   wait can be set by an appropriate choice of MTU.

   For example, if the MTU is set to 1500 bytes, then an MTU-sized
   packet will take about 8 milliseconds to send on a T1 (1.536 Mb/s)
   link.  But if the link speed is 19.2kb/s, then the transmission time
   becomes 625 ms -- well above our 100-200ms limit.  A 256-byte MTU
   would lower this delay to a little over 100 ms.  However, care should
   be taken not to lower the MTU excessively, as this will increase
   header overhead and trigger frequent router fragmentation (if Path
   MTU discovery is not in use).  This is likely to be the case with
   multicast, where Path MTU discovery is ineffective.

   One way to limit delay for interactive traffic without imposing a
   small MTU is to give priority to this traffic and to preempt (abort)



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   the transmission of a lower-priority packet when a higher priority
   packet arrives in the queue.  However, the link resources used to
   send the aborted packet are lost, and overall throughput will
   decrease.

   Another way to limit delay is to implement a link-level multiplexing
   scheme that allows several packets to be in progress simultaneously,
   with transmission priority given to segments of higher-priority IP
   packets.  For links using the Point-To-Point Protocol (PPP)
   [RFC1661], multi-class multilink [RFC2686] [RFC2687] [RFC2689]
   provides such a facility.

   ATM (asynchronous transfer mode), where SNDUs are fragmented and
   interleaved across smaller 53-byte ATM cells, is another example of
   this technique.  However, ATM is generally used on high-speed links
   where the store-and-forward delays are already minimal, and it
   introduces significant (~9%) increases in overhead due to the
   addition of 5-byte cell overhead to each 48-byte ATM cell.

   A third example is the Data-Over-Cable Service Interface
   Specification (DOCSIS) with typical upstream bandwidths of 2.56 Mb/s
   or 5.12 Mb/s.  To reduce the impact of a 1500-byte MTU in DOCSIS 1.0
   [DOCSIS1], a data link layer fragmentation mechanism is specified in
   DOCSIS 1.1 [DOCSIS2].  To accommodate the installed base, DOCSIS 1.1
   must be backward compatible with DOCSIS 1.0 cable modems, which
   generally do not support fragmentation.  Under the co-existence of
   DOCSIS 1.0 and DOCSIS 1.1, the unfragmented large data packets from
   DOCSIS 1.0 cable modems may affect the quality of service for voice
   packets from DOCSIS 1.1 cable modems.  In this case, it has been
   shown in [DOCSIS3] that the use of bandwidth allocation algorithms
   can mitigate this effect.

   To summarize, there is a fundamental tradeoff between efficiency and
   latency in the design of a subnetwork, and the designer should keep
   this tradeoff in mind.

3.  Framing on Connection-Oriented Subnetworks

   IP requires that subnetworks mark the beginning and end of each
   variable-length, asynchronous IP packet.  Some examples of links and
   subnetworks that do not provide this as an intrinsic feature include:

   1.  leased lines carrying a synchronous bit stream;

   2.  ISDN B-channels carrying a synchronous octet stream;

   3.  dialup telephone modems carrying an asynchronous octet stream;




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       and

   4.  Asynchronous Transfer Mode (ATM) networks carrying an
       asynchronous stream of fixed-sized "cells".

   The Internet community has defined packet framing methods for all
   these subnetworks.  The Point-To-Point Protocol (PPP) [RFC1661],
   which uses a variant of HDLC, is applicable to bit synchronous,
   octet-synchronous, and octet asynchronous links (i.e., examples 1-3
   above).  PPP is one preferred framing method for IP, since a large
   number of systems interoperate with PPP.  ATM has its own framing
   methods, described in [RFC2684] [RFC2364].

   At high speeds, a subnetwork should provide a framed interface
   capable of carrying asynchronous, variable-length IP datagrams.  The
   maximum packet size supported by this interface is discussed above in
   the MTU/Fragmentation section.  The subnetwork may implement this
   facility in any convenient manner.

   IP packet boundaries need not coincide with any framing or
   synchronization mechanisms internal to the subnetwork.  When the
   subnetwork implements variable sized data units, the most
   straightforward approach is to place exactly one IP packet into each
   subnetwork data unit (SNDU), and to rely on the subnetwork's existing
   ability to delimit SNDUs to also delimit IP packets.  A good example
   is Ethernet.  However, some subnetworks have SNDUs of one or more
   fixed sizes, as dictated by switching, forward error correction
   and/or interleaving considerations.  Examples of such subnetworks
   include ATM, with a single cell payload size of 48 octets plus a 5-
   octet header, and IS-95 digital cellular, with two "rate sets" of
   four fixed frame sizes each that may be selected on 20 millisecond
   boundaries.

   Because IP packets are of variable length, they may not necessarily
   fit into an integer multiple of fixed-sized SNDUs.  An "adaptation
   layer" is needed to convert IP packets into SNDUs while marking the
   boundary between each IP packet in some manner.

   There are several approaches to this problem.  The first is to encode
   each IP packet into one or more SNDUs with no SNDU containing pieces
   of more than one IP packet, and to pad out the last SNDU of the
   packet as needed.  Bits in a control header added to each SNDU
   indicate where the data segment belongs in the IP packet.  If the
   subnetwork provides in-order, at-most-once delivery, the header can
   be as simple as a pair of bits indicating whether the SNDU is the
   first and/or the last in the IP packet.  Alternatively, for
   subnetworks that do not reorder the fragments of an SNDU, only the
   last SNDU of the packet could be marked, as this would implicitly



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   indicate the next SNDU as the first in a new IP packet.  The AAL5
   (ATM Adaptation Layer 5) scheme used with ATM is an example of this
   approach, though it adds other features, including a payload length
   field and a payload CRC.

   In AAL5, the ATM User-User Indication, which is encoded in the
   Payload Type field of an ATM cell, indicates the last cell of a
   packet.  The packet trailer is located at the end of the SNDU and
   contains the packet length and a CRC.

   Another framing technique is to insert per-segment overhead to
   indicate the presence of a segment option.  When present, the option
   carries a pointer to the end of the packet.  This differs from AAL5
   in that it permits another packet to follow within the same segment.
   MPEG-2 Transport Streams [EN301192] [ISO13818] support this style of
   fragmentation, and may either use padding (limiting each MPEG
   transport stream packet to carry only part of one IP packet), or
   allow a second IP packet to start in the same Transport Stream packet
   (no padding).

   A third approach is to insert a special flag sequence into the data
   stream between each IP packet, and to pack the resulting data stream
   into SNDUs without regard to SNDU boundaries.  This may have
   implications when frames are lost.  The flag sequence can also pad
   unused space at the end of an SNDU.  If the special flag appears in
   the user data, it is escaped to an alternate sequence (usually larger
   than a flag) to avoid being misinterpreted as a flag.  The HDLC-based
   framing schemes used in PPP are all examples of this approach.

   All three adaptation schemes introduce overhead; how much depends on
   the distribution of IP packet sizes, the size(s) of the SNDUs, and in
   the HDLC-like approaches, the content of the IP packet (since flag-
   like sequences occurring in the packet must be escaped, which expands
   them).  The designer must also weigh implementation complexity and
   performance in the choice and design of an adaptation layer.

4.  Connection-Oriented Subnetworks

   IP has no notion of a "connection"; it is a purely connectionless
   protocol.  When a connection is required by an application, it is
   usually provided by TCP [RFC793], the Transmission Control Protocol,
   running atop IP on an end-to-end basis.

   Connection-oriented subnetworks can be (and are widely) used to carry
   IP, but often with considerable complexity.  Subnetworks consisting
   of few nodes can simply open a permanent connection between each pair
   of nodes.  This is frequently done with ATM.  However, the number of
   connections increases as the square of the number of nodes, so this



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   is clearly impractical for large subnetworks.  A "shim" layer between
   IP and the subnetwork is therefore required to manage connections.
   This is one of the most common functions of a Subnetwork Dependent
   Convergence Function (SNDCF) sublayer between IP and a subnetwork.

   SNDCFs typically open subnetwork connections as needed when an IP
   packet is queued for transmission and close them after an idle
   timeout.  There is no relation between subnetwork connections and any
   connections that may exist at higher layers (e.g., TCP).

   Because Internet traffic is typically bursty and transaction-
   oriented, it is often difficult to pick an optimal idle timeout.  If
   the timeout is too short, subnetwork connections are opened and
   closed rapidly, possibly over-stressing the subnetwork connection
   management system (especially if it was designed for voice traffic
   call holding times).  If the timeout is too long, subnetwork
   connections are idle much of the time, wasting any resources
   dedicated to them by the subnetwork.

   Purely connectionless subnets (such as Ethernet), which have no state
   and dynamically share resources, are optimal for supporting best-
   effort IP, which is stateless and dynamically shares resources.
   Connection-oriented packet networks (such as ATM and Frame Relay),
   which have state and dynamically share resources, are less optimal,
   since best-effort IP does not benefit from the overhead of creating
   and maintaining state.  Connection-oriented circuit-switched networks
   (including the PSTN and ISDN) have state and statically allocate
   resources for a call, and thus require state creation and maintenance
   overhead, but do not benefit from the efficiencies of statistical
   multiplexing sharing of capacity inherent in IP.

   In any event, if an SNDCF that opens and closes subnet connections is
   used to support IP, care should be taken to make sure that connection
   processing in the subnet can keep up with relatively short holding
   times.

5.  Broadcasting and Discovery

   Subnetworks fall into two categories: point-to-point and shared.  A
   point-to-point subnet has exactly two endpoint components (hosts or
   routers); a shared link has more than two endpoint components, using
   either an inherently broadcast medium (e.g., Ethernet, radio) or a
   switching layer hidden from the network layer (e.g., switched
   Ethernet, Myrinet [MYR95], ATM).  Switched subnetworks handle
   broadcast by copying broadcast packets, providing each interface that
   supports one, or more, systems (hosts or routers) with a copy of each
   packet.




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   Several Internet protocols for IPv4 make use of broadcast
   capabilities, including link-layer address lookup (ARP), auto-
   configuration (RARP, BOOTP, DHCP), and routing (RIP).

   A lack of broadcast capability can impede the performance of these
   protocols, or render them inoperable (e.g., DHCP).  ARP-like link
   address lookup can be provided by a centralized database, but at the
   expense of potentially higher response latency and the need for nodes
   to have explicit knowledge of the ARP server address.  Shared links
   should support native, link-layer subnet broadcast.

   A corresponding set of IPv6 protocols uses multicasting (see next
   section) instead of broadcasting to provide similar functions with
   improved scaling in large networks.

6.  Multicasting

   The Internet model includes "multicasting", where IP packets are sent
   to all the members of a multicast group [RFC1112] [RFC3376]
   [RFC2710].  Multicast is an option in IPv4, but a standard feature of
   IPv6.  IPv4 multicast is currently used by multimedia,
   teleconferencing, gaming, and file distribution (web, peer-to-peer
   sharing) applications, as well as by some key network and host
   protocols (e.g., RIPv2, OSPF, NTP).  IPv6 additionally relies on
   multicast for network configuration (DHCP-like autoconfiguration) and
   link-layer address discovery [RFC2461] (replacing ARP).  In the case
   of IPv6, this can allow autoconfiguration and address discovery to
   span across routers, whereas the IPv4 broadcast-based services cannot
   without ad-hoc router support [RFC1812].

   Multicast-enabled IP routers organize each multicast group into a
   spanning tree, and route multicast packets by making copies of each
   multicast packet and forwarding the copies to each output interface
   that includes at least one downstream member of the multicast group.

   Multicasting is considerably more efficient when a subnetwork
   explicitly supports it.  For example, a router relaying a multicast
   packet onto an Ethernet segment need send only one copy of the
   packet, no matter how many members of the multicast group are
   connected to the segment.  Without native multicast support, routers
   and switches on shared links would need to use broadcast with
   software filters, such that every multicast packet sent incurs
   software overhead for every node on the subnetwork, even if a node is
   not a member of the multicast group.  Alternately, the router would
   transmit a separate copy to every member of the multicast group on
   the segment, as is done on multicast-incapable switched subnets.





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   Subnetworks using shared channels (e.g., radio LANs, Ethernets) are
   especially suitable for native multicasting, and their designers
   should make every effort to support it.  This involves designating a
   section of the subnetwork's own address space for multicasting.  On
   these networks, multicast is basically broadcast on the medium, with
   Layer-2 receiver filters.

   Subnet interfaces also need to be designed to accept packets
   addressed to some number of multicast addresses, in addition to the
   unicast packets specifically addressed to them.  The number of
   multicast addresses that needs to be supported by a host depends on
   the requirements of the associated host; at least several dozen will
   meet most current needs.

   On low-speed networks, the multicast address recognition function may
   be readily implemented in host software, but on high-speed networks,
   it should be implemented in subnetwork hardware.  This hardware need
   not be complete; for example, many Ethernet interfaces implement a
   "hashing" function where the IP layer receives all of the multicast
   (and unicast) traffic to which the associated host subscribes, plus
   some small fraction of multicast traffic to which the host does not
   subscribe.  Host/router software then has to discard the unwanted
   packets that pass the Layer-2 multicast address filter [RFC1112].

   There does not need to be a one-to-one mapping between a Layer-2
   multicast address and an IP multicast address.  An address overlap
   may significantly degrade the filtering capability of a receiver's
   hardware multicast address filter.  A subnetwork supporting only
   broadcast should use this service for multicast and must rely on
   software filtering.

   Switched subnetworks must also provide a mechanism for copying
   multicast packets to ensure the packets reach at least all members of
   a multicast group.  One option is to "flood" multicast packets in the
   same manner as broadcast.  This can lead to unnecessary transmissions
   on some subnetwork links (notably non-multicast-aware Ethernet
   switches).  Some subnetworks therefore allow multicast filter tables
   to control which links receive packets belonging to a specific group.
   To configure this automatically requires access to Layer-3 group
   membership information (e.g., IGMP [RFC3376], or MLD [RFC2710]).
   Various implementation options currently exist to provide a subnet
   node with a list of mappings of multicast addresses to
   ports/interfaces.  These employ a range of approaches, including
   signaling from end hosts (e.g., IEEE 802 GARP/GMRP [802.1p]),
   signaling from switches (e.g., CGMP [CGMP] and RGMP [RFC3488]),
   interception and proxy of IP group membership packets (e.g., IGMP/MLD
   Proxy [MAGMA-PROXY]), and enabling Layer-2 devices to
   snoop/inspect/peek into forwarded Layer-3 protocol headers (e.g.,



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   IGMP, MLD, PIM) so that they may infer Layer-3 multicast group
   membership [MAGMA-SNOOP].  These approaches differ in their
   complexity, flexibility, and ability to support new protocols.

7.  Bandwidth on Demand (BoD) Subnets

   Some subnets allow a number of subnet nodes to share a channel
   efficiently by assigning transmission opportunities dynamically.
   Transmission opportunities are requested by a subnet node when it has
   packets to send.  The subnet schedules and grants transmission
   opportunities sufficient to allow the transmitting subnet node to
   send one or more packets (or packet fragments).  We call these
   subnets Bandwidth on Demand (BoD) subnets.  Examples of BoD subnets
   include Demand Assignment Multiple Access (DAMA) satellite and
   terrestrial wireless networks, IEEE 802.11 point coordination
   function (PCF) mode, and DOCSIS.  A connection-oriented network (such
   as the PSTN, ATM or Frame Relay) reserves resources on a much longer
   timescale, and is therefore not a BoD subnet in our taxonomy.

   The design parameters for BoD are similar to those in connection-
   oriented subnetworks, although the implementations may vary
   significantly.  In BoD, the user typically requests access to the
   shared channel for some duration.  Access may be allocated for a
   period of time at a specific rate, for a certain number of packets,
   or until the user releases the channel.  Access may be coordinated
   through a central management entity or with a distributed algorithm
   amongst the users.  Examples of the resource that may be shared
   include a terrestrial wireless hop, an upstream channel in a cable
   television system, a satellite uplink, and an end-to-end satellite
   channel.

   Long-delay BoD subnets pose problems similar to connection-oriented
   subnets in anticipating traffic.  While connection-oriented subnets
   hold idle channels open expecting new data to arrive, BoD subnets
   request channel access based on buffer occupancy (or expected buffer
   occupancy) on the sending port.  Poor performance will likely result
   if the sender does not anticipate additional traffic arriving at that
   port during the time it takes to grant a transmission request.  It is
   recommended that the algorithm have the capability to extend a hold
   on the channel for data that has arrived after the original request
   was generated (this may be done by piggybacking new requests on user
   data).

   There is a wide variety of BoD protocols available.  However, there
   has been relatively little comprehensive research on the interactions
   between BoD mechanisms and Internet protocol performance.  Research
   on some specific mechanisms is available (e.g., [AR02]).  One item
   that has been studied is TCP's retransmission timer [KY02].  BoD



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   systems can cause spurious timeouts when adjusting from a relatively
   high data rate, to a relatively low data rate.  In this case, TCP's
   transmitted data takes longer to get through the network than
   predicted by the TCP sender's computed retransmission timeout.
   Therefore, the TCP sender is prone to resending a segment
   prematurely.

8.  Reliability and Error Control

   In the Internet architecture, the ultimate responsibility for error
   recovery is at the end points [SRC81].  The Internet may occasionally
   drop, corrupt, duplicate, or reorder packets, and the transport
   protocol (e.g., TCP) or application (e.g., if UDP is used as the
   transport protocol) must recover from these errors on an end-to-end
   basis [RFC3155].  Error recovery in the subnetwork is therefore
   justifiable only to the extent that it can enhance overall
   performance.  It is important to recognize that a subnetwork can go
   too far in attempting to provide error recovery services in the
   Internet environment.  Subnet reliability should be "lightweight",
   i.e., it only has to be "good enough", *not* perfect.

   In this section, we discuss how to analyze characteristics of a
   subnetwork to determine what is "good enough".  The discussion below
   focuses on TCP, which is the most widely-used transport protocol in
   the Internet.  It is widely believed (and is a stated goal within the
   IETF) that non-TCP transport protocols should attempt to be "TCP-
   friendly" and have many of the same performance characteristics.
   Thus, the discussion below should be applicable, even to portions of
   the Internet where TCP may not be the predominant protocol.

8.1.  TCP vs Link-Layer Retransmission

   Error recovery involves the generation and transmission of redundant
   information computed from user data.  Depending on how much redundant
   information is sent and how it is generated, the receiver can use it
   to reliably detect transmission errors, correct up to some maximum
   number of transmission errors, or both.  The general approach is
   known as Error Control Coding, or ECC.

   The use of ECC to detect transmission errors so that retransmissions
   (hopefully without errors) can be requested is widely known as "ARQ"
   (Automatic Repeat Request).

   When enough ECC information is available to permit the receiver to
   correct some transmission errors without a retransmission, the
   approach is known as Forward Error Correction (FEC).  Due to the
   greater complexity of the required ECC and the need to tailor its
   design to the characteristics of a specific modem and channel, FEC



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   has traditionally been implemented in special-purpose hardware
   integral to a modem.  This effectively makes it part of the physical
   layer.

   Unlike ARQ, FEC was rarely used for telecommunications outside of
   space links prior to the 1990s.  It is now nearly universal in
   telephone, cable and DSL modems, digital satellite links, and digital
   mobile telephones.  FEC is also heavily used in optical and magnetic
   storage where "retransmissions" are not possible.

   Some systems use hybrid combinations of ARQ layered atop FEC; V.90
   dialup modems (in the upstream direction) with V.42 error control are
   one example.  Most errors are corrected by the trellis (FEC) code
   within the V.90 modem, and most remaining errors are detected and
   corrected by the ARQ mechanisms in V.42.

   Work is now underway to apply FEC above the physical layer, primarily
   in connection with reliable multicasting [RFC3048] [RFC3450-RFC3453]
   where conventional ARQ mechanisms are inefficient or difficult to
   implement.  However, in this discussion, we will assume that if FEC
   is present, it is implemented within the physical layer.

   Depending on the layer in which it is implemented, error control can
   operate on an end-to-end basis or over a shorter span, such as a
   single link.  TCP is the most important example of an end-to-end
   protocol that uses an ARQ strategy.

   Many link-layer protocols use ARQ, usually some flavor of HDLC
   [ISO3309].  Examples include the X.25 link layer, the AX.25 protocol
   used in amateur packet radio, 802.11 wireless LANs, and the reliable
   link layer specified in IEEE 802.2.

   Only end-to-end error recovery can ensure reliable service to the
   application (see Section 8).  However, some subnetworks (e.g., many
   wireless links) also have link-layer error recovery as a performance
   enhancement [RFC3366].  For example, many cellular links have small
   physical frame sizes (< 100 bytes) and relatively high frame loss
   rates.  Relying solely on end-to-end error recovery can clearly yield
   a performance degradation, as retransmissions across the end-to-end
   path take much longer to be received than when link layer
   retransmissions are used.  Thus, link-layer error recovery can often
   increase end-to-end performance.  As a result, link-layer and end-
   to-end recovery often co-exist; this can lead to the possibility of
   inefficient interactions between the two layers of ARQ protocols.

   This inter-layer "competition" might lead to the following wasteful
   situation.  When the link layer retransmits (parts of) a packet, the
   link latency momentarily increases.  Since TCP bases its



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   retransmission timeout on prior measurements of total end-to-end
   latency, including that of the link in question, this sudden increase
   in latency may trigger an unnecessary retransmission by TCP of a
   packet that the link layer is still retransmitting.  Such spurious
   end-to-end retransmissions generate unnecessary load and reduce end-
   to-end throughput.  As a result, the link layer may even have
   multiple copies of the same packet in the same link queue at the same
   time.  In general, one could say the competing error recovery is
   caused by an inner control loop (link-layer error recovery) reacting
   to the same signal as an outer control loop (end-to-end error
   recovery) without any coordination between the loops.  Note that this
   is solely an efficiency issue; TCP continues to provide reliable
   end-to-end delivery over such links.

   This raises the question of how persistent a link-layer sender should
   be in performing retransmission [RFC3366].  We define the link-layer
   (LL) ARQ persistency as the maximum time that a particular link will
   spend trying to transfer a packet before it can be discarded.  This
   deliberately simplified definition says nothing about the maximum
   number of retransmissions, retransmission strategies, queue sizes,
   queuing disciplines, transmission delays, or the like.  The reason we
   use the term LL ARQ persistency, instead of a term such as "maximum
   link-layer packet holding time," is that the definition closely
   relates to link-layer error recovery.  For example, on links that
   implement straightforward error recovery strategies, LL ARQ
   persistency will often correspond to a maximum number of
   retransmissions permitted per link-layer frame.

   For link layers that do not or cannot differentiate between flows
   (e.g., due to network layer encryption), the LL ARQ persistency
   should be small.  This avoids any harmful effects or performance
   degradation resulting from indiscriminate high persistence.  A
   detailed discussion of these issues is provided in [RFC3366].

   However, when a link layer can identify individual flows and apply
   ARQ selectively [LKJK02], then the link ARQ persistency should be
   high for a flow using reliable unicast transport protocols (e.g.,
   TCP) and must be low for all other flows.  Setting the link ARQ
   persistency larger than the largest link outage allows TCP to rapidly
   restore transmission without needing to wait for a retransmission
   time out.  This generally improves TCP performance in the face of
   transient outages.  However, excessively high persistence may be
   disadvantageous; a practical upper limit of 30-60 seconds may be
   desirable.  Implementation of such schemes remains a research issue.
   (See also the following section "Recovery from Subnetwork Outages").






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   Many subnetwork designers have opportunities to reduce the
   probability of packet loss, e.g., with FEC, ARQ, and interleaving, at
   the cost of increased delay.  TCP performance improves with
   decreasing loss but worsens with increasing end-to-end delay, so it
   is important to find the proper balance through analysis and
   simulation.

8.2.  Recovery from Subnetwork Outages

   Some types of subnetworks, particularly mobile radio, are subject to
   frequent temporary outages.  For example, an active cellular data
   user may drive or walk into an area (such as a tunnel) that is out of
   range of any base station.  No packets will be delivered successfully
   until the user returns to an area with coverage.

   The Internet protocols currently provide no standard way for a
   subnetwork to explicitly notify an upper layer protocol (e.g., TCP)
   that it is experiencing an outage rather than severe congestion.

   Under these circumstances TCP will, after each unsuccessful
   retransmission, wait even longer before trying again; this is its
   "exponential back-off" algorithm.  Furthermore, TCP will not discover
   that the subnetwork outage has ended until its next retransmission
   attempt.  If TCP has backed off, this may take some time.  This can
   lead to extremely poor TCP performance over such subnetworks.

   It is therefore highly desirable that a subnetwork subject to outages
   does not silently discard packets during an outage.  Ideally, the
   subnetwork should define an interface to the next higher layer (i.e.,
   IP) that allows it to refuse packets during an outage, and to
   automatically ask IP for new packets when it is again able to deliver
   them.  If it cannot do this, then the subnetwork should hold onto at
   least some of the packets it accepts during an outage and attempt to
   deliver them when the outage ends.  When packets are discarded, IP
   should be notified so that the appropriate ICMP messages can be sent.

   Note that it is *not* necessary to completely avoid dropping packets
   during an outage.  The purpose of holding onto a packet during an
   outage, either in the subnetwork or at the IP layer, is so that its
   eventual delivery will implicitly notify TCP that the subnetwork is
   again operational.  This is to enhance performance, not to ensure
   reliability -- reliability, as discussed earlier, can only be ensured
   on an end-to-end basis.

   Only a few packets per TCP connection, including ACKs, need be held
   in this way to cause the TCP sender to recover from the additional
   losses once the flow resumes [RFC3366].




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   Because it would be a layering violation (and possibly a performance
   hit) for IP or a subnetwork layer to look at TCP headers (which would
   in any event be impossible if IPsec encryption [RFC2401] is in use),
   it would be reasonable for the IP or subnetwork layers to choose, as
   a design parameter, some small number of packets that will be
   retained during an outage.

8.3.  CRCs, Checksums and Error Detection

   The TCP [RFC793], UDP [RFC768], ICMP, and IPv4 [RFC791] protocols all
   use the same simple 16-bit 1's complement checksum algorithm
   [RFC1071] to detect corrupted packets.  The IPv4 header checksum
   protects only the IPv4 header, while the TCP, ICMP, and UDP checksums
   provide end-to-end error detection for both the transport pseudo
   header (including network and transport layer information) and the
   transport payload data.  Protection of the data is optional for
   applications using UDP [RFC768] for IPv4, but is required for IPv6.

   The Internet checksum is not very strong from a coding theory
   standpoint, but it is easy to compute in software, and various
   proposals to replace the Internet checksums with stronger checksums
   have failed.  However, it is known that undetected errors can and do
   occur in packets received by end hosts [SP2000].

   To reduce processing costs, IPv6 has no IP header checksum.  The
   destination host detects "important" errors in the IP header, such as
   the delivery of the packet to the wrong destination.  This is done by
   including the IP source and destination addresses (pseudo header) in
   the computation of the checksum in the TCP or UDP header, a practice
   already performed in IPv4.  Errors in other IPv6 header fields may go
   undetected within the network; this was considered a reasonable price
   to pay for a considerable reduction in the processing required by
   each router, and it was assumed that subnetworks would use a strong
   link CRC.

   One way to provide additional protection for an IPv4 or IPv6 header
   is by the authentication and packet integrity services of the IP
   Security (IPsec) protocol [RFC2401].  However, this may not be a
   choice available to the subnetwork designer.

   Most subnetworks implement error detection just above the physical
   layer.  Packets corrupted in transmission are detected and discarded
   before delivery to the IP layer.  A 16-bit cyclic redundancy check
   (CRC) is usually the minimum for error detection.  This is
   significantly more robust against most patterns of errors than the
   16-bit Internet checksum.  Note that the error detection properties
   of a specific CRC code diminish with increasing frame size.  The
   Point-to-Point Protocol [RFC1662] requires support of a 16-bit CRC



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   for each link frame, with a 32-bit CRC as an option.  (PPP is often
   used in conjunction with a dialup modem, which provides its own error
   control).  Other subnetworks, including 802.3/Ethernet, AAL5/ATM,
   FDDI, Token Ring, and PPP over SONET/SDH all use a 32-bit CRC.  Many
   subnetworks can also use other mechanisms to enhance the error
   detection capability of the link CRC (e.g., FEC in dialup modems,
   mobile radio and satellite channels).

   Any new subnetwork designed to carry IP should therefore provide
   error detection for each IP packet that is at least as strong as the
   32-bit CRC specified in [ISO3309].  While this will achieve a very
   low undetected packet error rate due to transmission errors, it will
   not (and need not) achieve a very low packet loss rate as the
   Internet protocols are better suited to dealing with lost packets
   than to dealing with corrupted packets [SRC81].

   Packet corruption may be, and is, also caused by bugs in host and
   router hardware and software.  Even if every subnetwork implemented
   strong error detection, it is still essential that end-to-end
   checksums are used at the receiving end host [SP2000].

   Designers of complex subnetworks consisting of internal links and
   packet switches should consider implementing error detection on an
   edge-to-edge basis to cover an entire SNDU (or IP packet).  A CRC
   would be generated at the entry point to the subnetwork and checked
   at the exit endpoint.  This may be used instead of, or in combination
   with, error detection at the interface to each physical link.  An
   edge-to-edge check has the significant advantage of protecting
   against errors introduced anywhere within the subnetwork, not just
   within its transmission links.  Examples of this approach include the
   way in which the Ethernet CRC-32 is handled by LAN bridges [802.1D].
   ATM AAL5 [ITU-I363] also uses an edge-to-edge CRC-32.

   Some specific applications may be tolerant of residual errors in the
   data they exchange, but removal of the link CRC may expose the
   network to an undesirable increase in undetected errors in the IP and
   transport headers.  Applications may also require a high level of
   error protection for control information exchanged by protocols
   acting above the transport layer.  One example is a voice codec,
   which is robust against bit errors in the speech samples.  For such
   mechanisms to work, the receiving application must be able to
   tolerate receiving corrupted data.  This also requires that an
   application uses a mechanism to signal that payload corruption is
   permitted and to indicate the coverage (headers and data) required to
   be protected by the subnetwork CRC.  The UDP-Lite protocol [RFC3828]
   is the first Internet standards track transport protocol supporting
   partial payload protection.  Receipt of corrupt data by arbitrary




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   application protocols carries a serious danger that a subnet delivers
   data with errors that remain undetected by the application and hence
   corrupt the communicated data [SRC81].

8.4.  How TCP Works

   One of TCP's functions is end-host based congestion control for the
   Internet.  This is a critical part of the overall stability of the
   Internet, so it is important that link-layer designers understand
   TCP's congestion control algorithms.

   TCP assumes that, at the most abstract level, the network consists of
   links and queues.  Queues provide output-buffering on links that are
   momentarily oversubscribed.  They smooth instantaneous traffic bursts
   to fit the link bandwidth.  When demand exceeds link capacity long
   enough to fill the queue, packets must be dropped.  The traditional
   action of dropping the most recent packet ("tail dropping") is no
   longer recommended [RFC2309] [RFC2914], but it is still widely
   practiced.

   TCP uses sequence numbering and acknowledgments (ACKs) on an
   end-to-end basis to provide reliable, sequenced delivery.  TCP ACKs
   are cumulative, i.e., each implicitly ACKs every segment received so
   far.  If a packet with an unexpected sequence number is received, the
   ACK field in the packets returned by the receiver will cease to
   advance.  Using an optional enhancement, TCP can send selective
   acknowledgments (SACKs) [RFC2018] to indicate which segments have
   arrived at the receiver.

   Since the most common cause of packet loss is congestion, TCP treats
   packet loss as an indication of potential Internet congestion along
   the path between TCP end hosts.  This happens automatically, and the
   subnetwork need not know anything about IP or TCP.  A subnetwork node
   simply drops packets whenever it must, though some packet-dropping
   strategies (e.g., RED) are more fair to competing flows than others.

   TCP recovers from packet losses in two different ways.  The most
   important mechanism is the retransmission timeout.  If an ACK fails
   to arrive after a certain period of time, TCP retransmits the oldest
   unacked packet.  Taking this as a hint that the network is congested,
   TCP waits for the retransmission to be ACKed before it continues, and
   it gradually increases the number of packets in flight as long as a
   timeout does not occur again.

   A retransmission timeout can impose a significant performance
   penalty, as the sender is idle during the timeout interval and
   restarts with a congestion window of one TCP segment following the




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   timeout.  To allow faster recovery from the occasional lost packet in
   a bulk transfer, an alternate scheme, known as "fast recovery", was
   introduced [RFC2581] [RFC2582] [RFC2914] [TCPF98].

   Fast recovery relies on the fact that when a single packet is lost in
   a bulk transfer, the receiver continues to return ACKs to subsequent
   data packets that do not actually acknowledge any newly-received
   data.  These are known as "duplicate acknowledgments" or "dupacks".
   The sending TCP can use dupacks as a hint that a packet has been lost
   and retransmit it without waiting for a timeout.  Dupacks effectively
   constitute a negative acknowledgment (NAK) for the packet sequence
   number in the acknowledgment field.  TCP waits until a certain number
   of dupacks (currently 3) are seen prior to assuming a loss has
   occurred; this helps avoid an unnecessary retransmission during
   out-of-sequence delivery.

   A technique called "Explicit Congestion Notification" (ECN) [RFC3168]
   allows routers to directly signal congestion to hosts without
   dropping packets.  This is done by setting a bit in the IP header.
   Since ECN support is likely to remain optional, the lack of an ECN
   bit must *never* be interpreted as a lack of congestion.  Thus, for
   the foreseeable future, TCP must interpret a lost packet as a signal
   of congestion.

   The TCP "congestion avoidance" [RFC2581] algorithm maintains a
   congestion window (cwnd) controlling the amount of data TCP may have
   in flight at any moment.  Reducing cwnd reduces the overall bandwidth
   obtained by the connection; similarly, raising cwnd increases
   performance, up to the limit of the available capacity.

   TCP probes for available network capacity by initially setting cwnd
   to one or two packets and then increasing cwnd by one packet for each
   ACK returned from the receiver.  This is TCP's "slow start"
   mechanism.  When a packet loss is detected (or congestion is signaled
   by other mechanisms), cwnd is reset to one and the slow start process
   is repeated until cwnd reaches one half of its previous setting
   before the reset.  Cwnd continues to increase past this point, but at
   a much slower rate than before.  If no further losses occur, cwnd
   will ultimately reach the window size advertised by the receiver.

   This is an "Additive Increase, Multiplicative Decrease" (AIMD)
   algorithm.  The steep decrease of cwnd in response to congestion
   provides for network stability; the AIMD algorithm also provides for
   fairness between long running TCP connections sharing the same path.







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8.5.  TCP Performance Characteristics

   Caveat

   Here we present a current "state-of-the-art" understanding of TCP
   performance.  This analysis attempts to characterize the performance
   of TCP connections over links of varying characteristics.

   Link designers may wish to use the techniques in this section to
   predict what performance TCP/IP may achieve over a new link-layer
   design.  Such analysis is encouraged.  Because this is a relatively
   new analysis, and the theory is based on single-stream TCP
   connections under "ideal" conditions, it should be recognized that
   the results of such analysis may differ from actual performance in
   the Internet.  That being said, we have done our best to provide the
   designers with helpful information to get an accurate picture of the
   capabilities and limitations of TCP under various conditions.

8.5.1.  The Formulae

   The performance of TCP's AIMD Congestion Avoidance algorithm has been
   extensively analyzed.  The current best formula for the performance
   of the specific algorithms used by Reno TCP (i.e., the TCP specified
   in [RFC2581]) is given by Padhye, et al. [PFTK98].  This formula is:

                                         MSS
           BW = --------------------------------------------------------
                RTT*sqrt(1.33*p) + RTO*p*[1+32*p^2]*min[1,3*sqrt(.75*p)]

   where

           BW   is the maximum TCP throughout achievable by an
                individual TCP flow
           MSS  is the TCP segment size being used by the connection
           RTT  is the end-to-end round trip time of the TCP connection
           RTO  is the packet timeout (based on RTT)
           p    is the packet loss rate for the path
                (i.e., .01 if there is 1% packet loss)

   Note that the speed of the links making up the Internet path does not
   explicitly appear in this formula.  Attempting to send faster than
   the slowest link in the path causes the queue to grow at the
   transmitter driving the bottleneck.  This increases the RTT, which in
   turn reduces the achievable throughput.

   This is currently considered to be the best approximate formula for
   Reno TCP performance.  A further simplification of this formula is
   generally made by assuming that RTO is approximately 5*RTT.



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   TCP is constantly being improved.  A simpler formula, which gives an
   upper bound on the performance of any AIMD algorithm which is likely
   to be implemented in TCP in the future, was derived by Ott, et al.
   [MSMO97].

                     MSS   1
           BW = C    --- -------
                     RTT sqrt(p)

   where C is 0.93.

8.5.2.  Assumptions

   Both formulae assume that the TCP Receiver Window is not limiting the
   performance of the connection.  Because the receiver window is
   entirely determined by end-hosts, we assume that hosts will maximize
   the announced receiver window to maximize their network performance.

   Both of these formulae allow BW to become infinite if there is no
   loss.  However, an Internet path will drop packets at bottlenecked
   queues if the load is too high.  Thus, a completely lossless TCP/IP
   network can never occur (unless the network is being underutilized).

   The RTT used is the arithmetic average, including queuing delays.

   The formulae are for a single TCP connection.  If a path carries many
   TCP connections, each will follow the formulae above independently.

   The formulae assume long-running TCP connections.  For connections
   that are extremely short (<10 packets) and don't lose any packets,
   performance is driven by the TCP slow-start algorithm.  For
   connections of medium length, where on average only a few segments
   are lost, single connection performance will actually be slightly
   better than given by the formulae above.

   The difference between the simple and complex formulae above is that
   the complex formula includes the effects of TCP retransmission
   timeouts.  For very low levels of packet loss (significantly less
   than 1%), timeouts are unlikely to occur, and the formulae lead to
   very similar results.  At higher packet losses (1% and above), the
   complex formula gives a more accurate estimate of performance (which
   will always be significantly lower than the result from the simple
   formula).

   Note that these formulae break down as p approaches 100%.






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8.5.3.  Analysis of Link-Layer Effects on TCP Performance

   Consider the following example:

   A designer invents a new wireless link layer which, on average, loses
   1% of IP packets.  The link layer supports packets of up to 1040
   bytes, and has a one-way delay of 20 msec.

   If this link were to be used on an Internet path with a round trip
   time greater than 80ms, the upper bound may be computed by:

   For MSS, use 1000 bytes to exclude the 40 bytes of minimum IPv4 and
   TCP headers.

   For RTT, use 120 msec (80 msec for the Internet part, plus 20 msec
   each way for the new wireless link).

   For p, use .01.  For C, assume 1.

   The simple formula gives:

      BW = (1000 * 8 bits) / (.120 sec * sqrt(.01)) = 666 kbit/sec

   The more complex formula gives:

      BW = 402.9 kbit/sec

   If this were a 2 Mb/s wireless LAN, the designers might be somewhat
   disappointed.

   Some observations on performance:

   1.  We have assumed that the packet losses on the link layer are
       interpreted as congestion by TCP.  This is a "fact of life" that
       must be accepted.

   2.  The equations for TCP performance are all expressed in terms of
       packet loss, but many subnetwork designers think in terms of
       bit-error ratio.  *If* channel bit errors are independent, then
       the probability of a packet being corrupted is:

         p = 1 - ([1 - BER]^[FRAME_SIZE*8])

       Here we assume FRAME_SIZE is in bytes and "^" represents
       exponentiation.  It includes the user data and all headers
       (TCP,IP and subnetwork).  (Note: this analysis assumes the





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       subnetwork does not perform ARQ or transparent fragmentation
       [RFC3366].)  If the inequality

         BER * [FRAME_SIZE*8] << 1

       holds, the packet loss probability p can be approximated by:

         p = BER * [FRAME_SIZE*8]

       These equations can be used to apply BER to the performance
       equations above.

       Note that FRAME_SIZE can vary from one packet to the next.  Small
       packets (such as TCP acks) generally have a smaller probability
       of packet error than, say, a TCP packet carrying one MSS (maximum
       segment size) of user data.  A flow of small TCP acks can be
       expected to be slightly more reliable than a stream of larger TCP
       data segments.

       It bears repeating that the above analysis assumes that bit
       errors are statistically independent.  Because this is not true
       for many real links, our computation of p is actually an upper
       bound, not the exact probability of packet loss.

       There are many reasons why bit errors are not independent on real
       links.  Many radio links are affected by propagation fading or by
       interference that lasts over many bit times.  Also, links with
       Forward Error Correction (FEC) generally have very non-uniform
       bit error distributions that depend on the type of FEC, but in
       general the uncorrected errors tend to occur in bursts even when
       channel symbol errors are independent.  In all such cases, our
       computation of p from BER can only place an upper limit on the
       packet loss rate.

       If the distribution of errors under the FEC scheme is known, one
       could apply the same type of analysis as above, using the correct
       distribution function for the BER.  It is more likely in these
       FEC cases, however, that empirical methods are needed to
       determine the actual packet loss rate.

   3.  Note that the packet size plays an important role.  If the
       subnetwork loss characteristics are such that large packets have
       the same probability of loss as smaller packets, then larger
       packets will yield improved performance.







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   4.  We have chosen a specific RTT that might occur on a wide-area
       Internet path within the USA.  It is important to recognize that
       a variety of RTT values are experienced in the Internet.

       For example, RTTs are typically less than 10 msec in a wired LAN
       environment when communicating with a local host.  International
       connections may have RTTs of 200 msec or more.  Modems and other
       low-capacity links can add considerable delay due to their long
       packet transmission (serialisation) times.

       Links over geostationary repeater satellites have one-way speed-
       of-light delays of around 250ms, a minimum of 125ms propagation
       delay up to the satellite and 125ms down.  The RTT of an end-to-
       end TCP connection that includes such a link can be expected to
       be greater than 250ms.

       Queues on heavily-congested links may back up, increasing RTTs.
       Finally, virtual private networks (VPNs) and other forms of
       encryption and tunneling can add significant end-to-end delay to
       network connections.

9.  Quality-of-Service (QoS) considerations

   It is generally recognized that specific service guarantees are
   needed to support real-time multimedia, toll-quality telephony, and
   other performance-critical applications.  The provision of such
   Quality of Service guarantees in the Internet is an active area of
   research and standardization.  The IETF has not converged on a single
   service model, set of services, or single mechanism that will offer
   useful guarantees to applications and be scalable to the Internet.
   Indeed, the IETF does not have a single definition of Quality of
   Service.  [RFC2990] represents a current understanding of the
   challenges in architecting QoS for the Internet.

   There are presently two architectural approaches to providing
   mechanisms for QoS support in the Internet.

   IP Integrated Services (Intserv) [RFC1633] provides fine-grained
   service guarantees to individual flows.  Flows are identified by a
   flow specification (flowspec), which creates a stateful association
   between individual packets by matching fields in the packet header.
   Capacity is reserved for the flow, and appropriate traffic
   conditioning and scheduling is installed in routers along the path.
   The ReSerVation Protocol (RSVP) [RFC2205] [RFC2210] is usually, but
   need not necessarily be, used to install the flow QoS state.  Intserv
   defines two services, in addition to the Default (best effort)
   service.




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   1.  Guaranteed Service (GS) [RFC2212] offers hard upper bounds on
       delay to flows that conform to a traffic specification (TSpec).
       It uses a fluid-flow model to relate the TSpec and reserved
       bandwidth (RSpec) to variable delay.  Non-conforming packets are
       forwarded on a best-effort basis.

   2.  Controlled Load Service (CLS) [RFC2211] offers delay and packet
       loss equivalent to that of an unloaded network to flows that
       conform to a TSpec, but no hard bounds.  Non-conforming packets
       are forwarded on a best-effort basis.

   Intserv requires installation of state information in every
   participating router.  Performance guarantees cannot be made unless
   this state is present in every router along the path.  This, along
   with RSVP processing and the need for usage-based accounting, is
   believed to have scalability problems, particularly in the core of
   the Internet [RFC2208].

   IP Differentiated Services (Diffserv) [RFC2475] provides a "toolkit"
   offering coarse-grained controls to aggregates of flows.  Diffserv in
   itself does *not* provide QoS guarantees, but can be used to
   construct services with QoS guarantees across a Diffserv domain.
   Diffserv attempts to address the scaling issues associated with
   Intserv by requiring state awareness only at the edge of a Diffserv
   domain.  At the edge, packets are classified into flows, and the
   flows are conditioned (marked, policed, or shaped) to a traffic
   conditioning specification (TCS).  A Diffserv Codepoint (DSCP),
   identifying a per-hop behavior (PHB), is set in each packet header.
   The DSCP is carried in the DS-field, subsuming six bits of the former
   Type-of-Service (ToS) byte [RFC791] of the IP header [RFC2474].   The
   PHB denotes the forwarding behavior to be applied to the packet in
   each node in the Diffserv domain.  Although there is a "recommended"
   DSCP associated with each PHB, the mappings from DSCPs to PHBs are
   defined by the DS-domain.  In fact, there can be several DSCPs
   associated with the same PHB.  Diffserv presently defines three PHBs.

   1.  The class selector PHB [RFC2474] replaces the IP precedence field
       of the former ToS byte.  It offers relative forwarding
       priorities.

   2.  The Expedited Forwarding (EF) PHB [RFC3246] [RFC3248] guarantees
       that packets will have a well-defined minimum departure rate
       which, if not exceeded, ensures that the associated queues are
       short or empty.  EF is intended to support services that offer
       tightly-bounded loss, delay, and delay jitter.






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   3.  The Assured Forwarding (AF) PHB group [RFC2597] offers different
       levels of forwarding assurance for each aggregated flow of
       packets.  Each AF group is independently allocated forwarding
       resources.  Packets are marked with one of three drop
       precedences; those with the highest drop precedence are dropped
       with lower probability than those marked with the lowest drop
       precedence.  DSCPs are recommended for four independent AF
       groups, although a DS domain can have more or fewer AF groups.

   Ongoing work in the IETF is addressing ways to support Intserv with
   Diffserv.  There is some belief (e.g., as expressed in [RFC2990])
   that such an approach will allow individual flows to receive service
   guarantees and scale to the global Internet.

   The QoS guarantees that can be offered by the IP layer are a product
   of two factors:

   1.  the concatenation of the QoS guarantees offered by the subnets
       along the path of a flow.  This implies that a subnet may wish to
       offer multiple services (with different QoS guarantees) to the IP
       layer, which can then determine which flows use which subnet
       service.  To put it another way, forwarding behavior in the
       subnet needs to be "clued" by the forwarding behavior (service or
       PHB) at the IP layer, and

   2.  the operation of a set of cooperating mechanisms, such as
       bandwidth reservation and admission control, policy management,
       traffic classification, traffic conditioning (marking, policing
       and/or shaping), selective discard, queuing, and scheduling.
       Note that support for QoS in subnets may require similar
       mechanisms, especially when these subnets are general topology
       subnets (e.g., ATM, frame relay, or MPLS) or shared media
       subnets.

   Many subnetwork designers face inherent tradeoffs between delay,
   throughput, reliability, and cost.  Other subnetworks have parameters
   that manage bandwidth, internal connection state, and the like.
   Therefore, the following subnetwork capabilities may be desirable,
   although some might be trivial or moot if the subnet is a dedicated
   point-to-point link.

   1.  The subnetwork should have the ability to reserve bandwidth for a
       connection or flow and schedule packets accordingly.

   2.  Bandwidth reservations should be based on a one- or two-token
       bucket model, depending on whether the service is intended to
       support constant-rate or bursty traffic.




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   3.  If a connection or flow does not use its reserved bandwidth at a
       given time, the unused bandwidth should be available for other
       flows.

   4.  Packets in excess of a connection or flow's agreed rate should be
       forwarded as best-effort or discarded, depending on the service
       offered by the subnet to the IP layer.

   5.  If a subnet contains error control mechanisms (retransmission
       and/or FEC), it should be possible for the IP layer to influence
       the inherent tradeoffs between uncorrected errors, packet losses,
       and delay.  These capabilities at the subnet/IP layer service
       boundary correspond to selection of more or less error control
       and/or to selection of particular error control mechanisms within
       the subnetwork.

   6.  The subnet layer should know, and be able to inform the IP layer,
       how much fixed delay and delay jitter it offers for a flow or
       connection.  If the Intserv model is used, the delay jitter
       component may be best expressed in terms of the TSpec/RSpec model
       described in [RFC2212].

   7.  Support of the Diffserv class selectors [RFC2474] suggests that
       the subnet might consider mechanisms that support priorities.

10.  Fairness vs Performance

   Subnetwork designers should be aware of the tradeoffs between
   fairness and efficiency inherent in many transmission scheduling
   algorithms.  For example, many local area networks use contention
   protocols to resolve access to a shared transmission channel.  These
   protocols represent overhead.  While limiting the amount of data that
   a subnet node may transmit per contention cycle helps assure timely
   access to the channel for each subnet node, it also increases
   contention overhead per unit of data sent.

   In some mobile radio networks, capacity is limited by interference,
   which in turn depends on average transmitter power.  Some receivers
   may require considerably more transmitter power (generating more
   interference and consuming more channel capacity) than others.

   In each case, the scheduling algorithm designer must balance
   competing objectives: providing a fair share of capacity to each
   subnet node while maximizing the total capacity of the network.  One
   approach for balancing performance and fairness is outlined in
   [ES00].





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11.  Delay Characteristics

   The TCP sender bases its retransmission timeout (RTO) on measurements
   of the round trip delay experienced by previous packets.  This allows
   TCP to adapt automatically to the very wide range of delays found on
   the Internet.  The recommended algorithms are described in [RFC2988].
   Evaluations of TCP's retransmission timer can be found in [AP99] and
   [LS00].

   These algorithms model the delay along an Internet path as a
   normally-distributed random variable with a slowly-varying mean and
   standard deviation.  TCP estimates these two parameters by
   exponentially smoothing individual delay measurements, and it sets
   the RTO to the estimated mean delay plus some fixed number of
   standard deviations.  (The algorithm actually uses mean deviation as
   an approximation to standard deviation, because it is easier to
   compute.)

   The goal is to compute an RTO that is small enough to detect and
   recover from packet losses while minimizing unnecessary ("spurious")
   retransmissions when packets are unexpectedly delayed but not lost.
   Although these goals conflict, the algorithm works well when the
   delay variance along the Internet path is low, or the packet loss
   rate is low.

   If the path delay variance is high, TCP sets an RTO that is much
   larger than the mean of the measured delays.  If the packet loss rate
   is low, the large RTO is of little consequence, as timeouts occur
   only rarely.  Conversely, if the path delay variance is low, then TCP
   recovers quickly from lost packets; again, the algorithm works well.
   However, when delay variance and the packet loss rate are both high,
   these algorithms perform poorly, especially when the mean delay is
   also high.

   Because TCP uses returning acknowledgments as a "clock" to time the
   transmission of additional data, excessively high delays (even if the
   delay variance is low) also affect TCP's ability to fully utilize a
   high-speed transmission pipe.  It also slows the recovery of lost
   packets, even when delay variance is small.

   Subnetwork designers should therefore minimize all three parameters
   (delay, delay variance, and packet loss) as much as possible.

   In many subnetworks, these parameters are inherently in conflict.
   For example, on a mobile radio channel, the subnetwork designer can
   use retransmission (ARQ) and/or forward error correction (FEC) to
   trade off delay, delay variance, and packet loss in an effort to
   improve TCP performance.  While ARQ increases delay variance, FEC



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   does not.  However, FEC (especially when combined with interleaving)
   often increases mean delay, even on good channels where ARQ
   retransmissions are not needed and ARQ would not increase either the
   delay or the delay variance.

   The tradeoffs among these error control mechanisms and their
   interactions with TCP can be quite complex, and are the subject of
   much ongoing research.  We therefore recommend that subnetwork
   designers provide as much flexibility as possible in the
   implementation of these mechanisms, and provide access to them as
   discussed above in the section on Quality of Service.

12.  Bandwidth Asymmetries

   Some subnetworks may provide asymmetric bandwidth (or may cause TCP
   packet flows to experience asymmetry in the capacity) and the
   Internet protocol suite will generally still work fine.  However,
   there is a case when such a scenario reduces TCP performance.  Since
   TCP data segments are "clocked" out by returning acknowledgments, TCP
   senders are limited by the rate at which ACKs can be returned
   [BPK98].  Therefore, when the ratio of the available capacity of the
   Internet path carrying the data to the bandwidth of the return path
   of the acknowledgments is too large, the slow return of the ACKs
   directly impacts performance.  Since ACKs are generally smaller than
   data segments, TCP can tolerate some asymmetry, but as a general
   rule, designers of subnetworks should be aware that subnetworks with
   significant asymmetry can result in reduced performance, unless
   issues are taken to mitigate this [RFC3449].

   Several strategies have been identified for reducing the impact of
   asymmetry of the network path between two TCP end hosts, e.g.,
   [RFC3449].  These techniques attempt to reduce the number of ACKs
   transmitted over the return path (low bandwidth channel) by changes
   at the end host(s), and/or by modification of subnetwork packet
   forwarding.  While these solutions may mitigate the performance
   issues caused by asymmetric subnetworks, they do have associated cost
   and may have other implications.  A fuller discussion of strategies
   and their implications is provided in [RFC3449].

13.  Buffering, flow and congestion control

   Many subnets include multiple links with varying traffic demands and
   possibly different transmission speeds.  At each link there must be a
   queuing system, including buffering, scheduling, and a capability to
   discard excess subnet packets.  These queues may also be part of a
   subnet flow control or congestion control scheme.





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   For the purpose of this discussion, we talk about packets without
   regard to whether they refer to a complete IP packet or a subnetwork
   frame.  At each queue, a packet experiences a delay that depends on
   competing traffic and the scheduling discipline, and is subjected to
   a local discarding policy.

   Some subnets may have flow or congestion control mechanisms in
   addition to packet dropping.  Such mechanisms can operate on
   components in the subnet layer, such as schedulers, shapers, or
   discarders, and can affect the operation of IP forwarders at the
   edges of the subnet.  However, with the exception of Explicit
   Congestion Notification [RFC3168] (discussed below), IP has no way to
   pass explicit congestion or flow control signals to TCP.

   TCP traffic, especially aggregated TCP traffic, is bursty.  As a
   result, instantaneous queue depths can vary dramatically, even in
   nominally stable networks.  For optimal performance, packets should
   be dropped in a controlled fashion, not just when buffer space is
   unavailable.  How much buffer space should be supplied is still a
   matter of debate, but as a rule of thumb, each node should have
   enough buffering to hold one link_bandwidth*link_delay product's
   worth of data for each TCP connection sharing the link.

   This is often difficult to estimate, since it depends on parameters
   beyond the subnetwork's control or knowledge.  Internet nodes
   generally do not implement admission control policies, and cannot
   limit the number of TCP connections that use them.  In general, it is
   wise to err in favor of too much buffering rather than too little.
   It may also be useful for subnets to incorporate mechanisms that
   measure propagation delays to assist in buffer sizing calculations.

   There is a rough consensus in the research community that active
   queue management is important to improving fairness, link
   utilization, and throughput [RFC2309].  Although there are questions
   and concerns about the effectiveness of active queue management
   (e.g., [MBDL99]), it is widely considered an improvement over tail-
   drop discard policies.

   One form of active queue management is the Random Early Detection
   (RED) algorithm [RED93], a family of related algorithms.  In one
   version of RED, an exponentially-weighted moving average of the queue
   depth is maintained:

      When this average queue depth is between a maximum threshold
      max_th and a minimum threshold min_th, the probability of packets
      that are dropped is proportional to the amount by which the
      average queue depth exceeds min_th.




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      When this average queue depth is equal to max_th, the drop
      probability is equal to a configurable parameter max_p.

      When this average queue depth is greater than max_th, packets are
      always dropped.

   Numerous variants on RED appear in the literature, and there are
   other active queue management algorithms which claim various
   advantages over RED [GM02].

   With an active queue management algorithm, dropped packets become a
   feedback signal to trigger more appropriate congestion behavior by
   the TCPs in the end hosts.  Randomization of dropping tends to break
   up the observed tendency of TCP windows belonging to different TCP
   connections to become synchronized by correlated drops, and it also
   imposes a degree of fairness on those connections that implement TCP
   congestion avoidance properly.  Another important property of active
   queue management algorithms is that they attempt to keep average
   queue depths short while accommodating large short-term bursts.

   Since TCP neither knows nor cares whether congestive packet loss
   occurs at the IP layer or in a subnet, it may be advisable for
   subnets that perform queuing and discarding to consider implementing
   some form of active queue management.  This is especially true if
   large aggregates of TCP connections are likely to share the same
   queue.  However, active queue management may be less effective in the
   case of many queues carrying smaller aggregates of TCP connections,
   e.g., in an ATM switch that implements per-VC queuing.

   Note that the performance of active queue management algorithms is
   highly sensitive to settings of configurable parameters, and also to
   factors such as RTT [MBB00] [FB00].

   Some subnets, most notably ATM, perform segmentation and reassembly
   at the subnetwork edges.  Care should be taken here in designing
   discard policies.  If the subnet discards a fragment of an IP packet,
   then the remaining fragments become an unproductive load on the
   subnet that can markedly degrade end-to-end performance [RF95].
   Subnetworks should therefore attempt to discard these extra fragments
   whenever one of them must be discarded.  If the IP packet has already
   been partially forwarded when discarding becomes necessary, then
   every remaining fragment except the one marking the end of the IP
   packet should also be discarded.  For ATM subnets, this specifically
   means using Early Packet Discard and Partial Packet Discard [ATMFTM].

   Some subnets include flow control mechanisms that effectively require
   that the rate of traffic flows be shaped upon entry to the subnet.
   One example of such a subnet mechanism is in the ATM Available Bit



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   rate (ABR) service category [ATMFTM].  Such flow control mechanisms
   have the effect of making the subnet nearly lossless by pushing
   congestion into the IP routers at the edges of the subnet.  In such a
   case, adequate buffering and discard policies are needed in these
   routers to deal with a subnet that appears to have varying bandwidth.
   Whether there is a benefit in this kind of flow control is
   controversial; there are numerous simulation and analytical studies
   that go both ways.  It appears that some of the issues leading to
   such different results include sensitivity to ABR parameters, use of
   binary rather than explicit rate feedback, use (or not) of per-VC
   queuing, and the specific ATM switch algorithms selected for the
   study.  Anecdotally, some large networks that used IP over ABR to
   carry TCP traffic have claimed it to be successful, but have
   published no results.

   Another possible approach to flow control in the subnet would be to
   work with TCP Explicit Congestion Notification (ECN) semantics
   [RFC3168] through utilizing explicit congestion indicators in subnet
   frames.  Routers at the edges of the subnet, rather than shaping,
   would set the explicit congestion bit in those IP packets that are
   received in subnet frames that have an ECN indication.  Nodes in the
   subnet would need to implement an active queue management protocol
   that marks subnet frames instead of dropping them.

   ECN is currently a proposed standard, but it is not yet widely
   deployed.

14.  Compression

   Application data compression is a function that can usually be
   omitted in the subnetwork.  The endpoints typically have more CPU and
   memory resources to run a compression algorithm and a better
   understanding of what is being compressed.  End-to-end compression
   benefits every network element in the path, while subnetwork-layer
   compression, by definition, benefits only a single subnetwork.

   Data presented to the subnetwork layer may already be in a compressed
   format (e.g., a JPEG file), compressed at the application layer
   (e.g., the optional "gzip", "compress", and "deflate" compression in
   HTTP/1.1 [RFC2616]), or compressed at the IP layer (the IP Payload
   Compression Protocol [RFC3173] supports DEFLATE [RFC2394] and LZS
   [RFC2395]).  Compression at the subnetwork edges is of no benefit for
   any of these cases.

   The subnetwork may also process data that has been encrypted by the
   application (OpenPGP [RFC2440] or S/MIME [RFC2633]), just above TCP
   (SSL, TLS [RFC2246]), or just above IP (IPsec ESP [RFC2406]).




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   Ciphers generate high-entropy bit streams lacking any patterns that
   can be exploited by a compression algorithm.

   However, much data is still transmitted uncompressed over the
   Internet, so subnetwork compression may be beneficial.  Any
   subnetwork compression algorithm must not expand uncompressible data,
   e.g., data that has already been compressed or encrypted.

   We make a strong recommendation that subnetworks operating at low
   speed or with small MTUs compress IP and transport-level headers (TCP
   and UDP) using several header compression schemes developed within
   the IETF [RFC3150].  An uncompressed 40-byte TCP/IP header takes
   about 33 milliseconds to send at 9600 bps.  "VJ" TCP/IP header
   compression [RFC1144] compresses most headers to 3-5 bytes, reducing
   transmission time to several milliseconds on dialup modem links.
   This is especially beneficial for small, latency-sensitive packets in
   interactive sessions.

   Similarly, RTP compression schemes, such as CRTP [RFC2508] and ROHC
   [RFC3095], compress most IP/UDP/RTP headers to 1-4 bytes.  The
   resulting savings are especially significant when audio packets are
   kept small to minimize store-and-forward latency.

   Designers should consider the effect of the subnetwork error rate on
   the performance of header compression.  TCP ordinarily recovers from
   lost packets by retransmitting only those packets that were actually
   lost; packets arriving correctly after a packet loss are kept on a
   resequencing queue and do not need to be retransmitted.  In VJ TCP/IP
   [RFC1144] header compression, however, the receiver cannot explicitly
   notify a sender of data corruption and subsequent loss of
   synchronization between compressor and decompressor.  It relies
   instead on TCP retransmission to re-synchronize the decompressor.
   After a packet is lost, the decompressor must discard every
   subsequent packet, even if the subnetwork makes no further errors,
   until the sending TCP retransmits to re-synchronize the decompressor.
   This effect can substantially magnify the effect of subnetwork packet
   losses if the sending TCP window is large, as it will often be on a
   path with a large bandwidth*delay product [LRKOJ99].

   Alternate header compression schemes, such as those described in
   [RFC2507], include an explicit request for retransmission of an
   uncompressed packet to allow decompressor resynchronization without
   waiting for a TCP retransmission.  However, these schemes are not yet
   in widespread use.

   Both TCP header compression schemes do not compress widely-used TCP
   options such as selective acknowledgements (SACK).  Both fail to
   compress TCP traffic that makes use of explicit congestion



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   notification (ECN).  Work is under way in the IETF ROHC WG to address
   these shortcomings in a ROHC header compression scheme for TCP
   [RFC3095] [RFC3096].

   The subnetwork error rate also is important for RTP header
   compression.  CRTP uses delta encoding, so a packet loss on the link
   causes uncertainty about the subsequent packets, which often must be
   discarded until the decompressor has notified the compressor and the
   compressor has sent re-synchronizing information.  This typically
   takes slightly more than the end-to-end path round-trip time.  For
   links that combine significant error rates with latencies that
   require multiple packets to be in flight at a time, this leads to
   significant error propagation, i.e., subsequent losses caused by an
   initial loss.

   For links that are both high-latency (multiple packets in flight from
   a typical RTP stream) and error-prone, RTP ROHC provides a more
   robust way of RTP header compression, at a cost of higher complexity
   at the compressor and decompressor.  For example, within a talk
   spurt, only extended losses of (depending on the mode chosen) 12-64
   packets typically cause error propagation.

15.  Packet Reordering

   The Internet architecture does not guarantee that packets will arrive
   in the same order in which they were originally transmitted;
   transport protocols like TCP must take this into account.

   However, reordering does come at a cost with TCP as it is currently
   defined.  Because TCP returns a cumulative acknowledgment (ACK)
   indicating the last in-order segment that has arrived, out-of-order
   segments cause a TCP receiver to transmit a duplicate acknowledgment.
   When the TCP sender notices three duplicate acknowledgments, it
   assumes that a segment was dropped by the network and uses the fast
   retransmit algorithm [Jac90] [RFC2581] to resend the segment.  In
   addition, the congestion window is reduced by half, effectively
   halving TCP's sending rate.  If a subnetwork reorders segments
   significantly such that three duplicate ACKs are generated, the TCP
   sender needlessly reduces the congestion window and performance
   suffers.

   Packet reordering frequently occurs in parts of the Internet, and it
   seems to be difficult or impossible to eliminate [BPS99].  For this
   reason, research on improving TCP's behavior in the face of packet
   reordering [LK00] [BA02] has begun.






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   [BPS99] cites reasons why it may even be undesirable to eliminate
   reordering.  There are situations where average packet latency can be
   reduced, link efficiency can be increased, and/or reliability can be
   improved if reordering is permitted.  Examples include certain high
   speed switches within the Internet backbone and the parallel links
   used over many Internet paths for load splitting and redundancy.

   This suggests that subnetwork implementers should try to avoid packet
   reordering whenever possible, but not if doing so compromises
   efficiency, impairs reliability, or increases average packet delay.

   Note that every header compression scheme currently standardized for
   the Internet requires in-order packet delivery on the link between
   compressor and decompressor.  PPP is frequently used to carry
   compressed TCP/IP packets; since it was originally designed for
   point-to-point and dialup links, it is assumed to provide in-order
   delivery.  For this reason, subnetwork implementers who provide PPP
   interfaces to VPNs and other more complex subnetworks, must also
   maintain in-order delivery of PPP frames.

16.  Mobility

   Internet users are increasingly mobile.  Not only are many Internet
   nodes laptop computers, but pocket organizers and mobile embedded
   systems are also becoming nodes on the Internet.  These nodes may
   connect to many different access points on the Internet over time,
   and they expect this to be largely transparent to their activities.
   Except when they are not connected to the Internet at all, and for
   performance differences when they are connected, they expect that
   everything will "just work" regardless of their current Internet
   attachment point or local subnetwork technology.

   Changing a host's Internet attachment point involves one or more of
   the following steps.

   First, if use of the local subnetwork is restricted, the user's
   credentials must be verified and access granted.  There are many ways
   to do this.  A trivial example would be an "Internet cafe" that
   grants physical access to the subnetwork for a fee.  Subnetworks may
   implement technical access controls of their own; one example is IEEE
   802.11 Wireless Equivalent Privacy [IEEE80211].  It is common
   practice for both cellular telephone and Internet service providers
   (ISPs) to agree to serve one anothers' users; RADIUS [RFC2865] is the
   standard method for ISPs to exchange authorization information.

   Second, the host may have to be reconfigured with IP parameters
   appropriate for the local subnetwork.  This usually includes setting
   an IP address, default router, and domain name system (DNS) servers.



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   On multiple-access networks, the Dynamic Host Configuration Protocol
   (DHCP) [RFC2131] is almost universally used for this purpose.  On PPP
   links, these functions are performed by the IP Control Protocol
   (IPCP) [RFC1332].

   Third, traffic destined for the mobile host must be routed to its
   current location.  This roaming function is the most common meaning
   of the term "Internet mobility".

   Internet mobility can be provided at any of several layers in the
   Internet protocol stack, and there is ongoing debate as to which is
   the most appropriate and efficient.  Mobility is already a feature of
   certain application layer protocols; the Post Office Protocol (POP)
   [RFC1939] and the Internet Message Access Protocol (IMAP) [RFC3501]
   were created specifically to provide mobility in the receipt of
   electronic mail.

   Mobility can also be provided at the IP layer [RFC3344].  This
   mechanism provides greater transparency, viz., IP addresses that
   remain fixed as the nodes move, but at the cost of potentially
   significant network overhead and increased delay because of the sub-
   optimal network routing and tunneling involved.

   Some subnetworks may provide internal mobility, transparent to IP, as
   a feature of their own internal routing mechanisms.  To the extent
   that these simplify routing at the IP layer, reduce the need for
   mechanisms like Mobile IP, or exploit mechanisms unique to the
   subnetwork, this is generally desirable.  This is especially true
   when the subnetwork covers a relatively small geographic area and the
   users move rapidly between the attachment points within that area.
   Examples of internal mobility schemes include Ethernet switching and
   intra-system handoff in cellular telephony.

   However, if the subnetwork is physically large and connects to other
   parts of the Internet at multiple geographic points, care should be
   taken to optimize the wide-area routing of packets between nodes on
   the external Internet and nodes on the subnet.  This is generally
   done with "nearest exit" routing strategies.  Because a given
   subnetwork may be unaware of the actual physical location of a
   destination on another subnetwork, it simply routes packets bound for
   the other subnetwork to the nearest router between the two.  This
   implies some awareness of IP addressing and routing within the
   subnetwork.  The subnetwork may wish to use IP routing internally for
   wide area routing and restrict subnetwork-specific routing to
   constrained geographic areas where the effects of suboptimal routing
   are minimized.





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17.  Routing

   Subnetworks connecting more than two systems must provide their own
   internal Layer-2 forwarding mechanisms, either implicitly (e.g.,
   broadcast) or explicitly (e.g., switched).  Since routing is the
   major function of the Internet layer, the question naturally arises
   as to the interaction between routing at the Internet layer and
   routing in the subnet, and proper division of function between the
   two.

   Layer-2 subnetworks can be point-to-point, connecting two systems, or
   multipoint.  Multipoint subnetworks can be broadcast (e.g., shared
   media or emulated) or non-broadcast.  Generally, IP considers
   multipoint subnetworks as broadcast, with shared-medium Ethernet as
   the canonical (and historical) example, and point-to-point
   subnetworks as a degenerate case.  Non-broadcast subnetworks may
   require additional mechanisms, e.g., above IP at the routing layer
   [RFC2328].

   IP is ignorant of the topology of the subnetwork layer.  In
   particular, reconfiguration of subnetwork paths is not tracked by the
   IP layer.  IP is only affected by whether it can send/receive packets
   sent to the remotely connected systems via the subnetwork interface
   (i.e., the reachability from one router to another).  IP further
   considers that subnetworks are largely static -- that both their
   membership and existence are stable at routing timescales (tens of
   seconds); changes to these are considered re-provisioning, rather
   than routing.

   Routing functionality in a subnetwork is related to addressing in
   that subnetwork.  Resolution of addresses on subnetwork links is
   required for forwarding IP packets across links (e.g., ARP for IPv4,
   or ND for IPv6).  There is unlikely to be direct interaction between
   subnetwork routing and IP routing.  Where broadcast is provided or
   explicitly emulated, address resolution can be used directly; where
   not provided, the link layer routing may interface to a protocol for
   resolution, e.g., to the Next-Hop Resolution Protocol [RFC2322] to
   provide context-dependent address resolution capabilities.

   Subnetwork routing can either complement or compete with IP routing.
   It complements IP when a subnetwork encapsulates its internal
   routing, and where the effects of that routing are not visible at the
   IP layer.  However, if different paths in the subnetwork have
   characteristics that affect IP routing, it can affect or even inhibit
   the convergence of IP routing.






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   Routing protocols generally consider Layer-2 subnetworks, i.e., with
   subnet masks and no intermediate IP hops, to have uniform routing
   metrics to all members.  Routing can break when a link's
   characteristics do not match the routing metric, in this case, e.g.,
   when some member pairs have different path characteristics.  Consider
   a virtual Ethernet subnetwork that includes both nearby (sub-
   millisecond latency) and remote (100's of milliseconds away) systems.
   Presenting that group as a single subnetwork means that some routing
   protocols will assume that all pairs have the same delay, and that
   that delay is small.  Because this is not the case, the routing
   tables constructed may be suboptimal or may even fail to converge.

   When a subnetwork is used for transit between a set of routers, it
   conventionally provides the equivalent of a full mesh of point-to-
   point links.  Simplicity of the internal subnet structure can be used
   (e.g., via NHRP [RFC2332]) to reduce the size of address resolution
   tables, but routing exchanges will continue to reflect the full mesh
   they emulate.  In general, subnetworks should not be used as a
   transit among a set of routers where routing protocols would break if
   a full mesh of equivalent point-to-point links were used.

   Some subnetworks have special features that allow the use of more
   effective or responsive routing mechanisms that cannot be implemented
   in IP because of its need for generality.  One example is the self-
   learning bridge algorithm widely used in Ethernet networks.  Learning
   bridges perform Layer-2 subnetwork forwarding, avoiding the need for
   dynamic routing at each subnetwork hop.  Another is the "handoff"
   mechanism in cellular telephone networks, particularly the "soft
   handoff" scheme in IS-95 CDMA.

   Subnetworks that cover large geographic areas or include links of
   widely-varying capabilities should be avoided.  IP routing generally
   considers all multipoint subnets equivalent to a local, shared-medium
   link with uniform metrics between any pair of systems, and ignores
   internal subnetwork topology.  Where a subnetwork diverges from that
   assumption, it is the obligation of subnetwork designers to provide
   compensating mechanisms.  Not doing so can affect the scalability and
   convergence of IP routing, as noted above.

   The subnetwork designer who decides to implement internal routing
   should consider whether a custom routing algorithm is warranted, or
   if an existing Internet routing algorithm or protocol may suffice.
   The designer should consider whether this decision is to reduce the
   address resolution table size (possible, but with additional protocol
   support required), or is trying to reduce routing table complexity.
   The latter may be better achieved by partitioning the subnetwork,
   either physically or logically, and using network-layer protocols to
   support partitioning (e.g., AS's in BGP).  Protocols and routing



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   algorithms can be notoriously subtle, complex, and difficult to
   implement correctly.  Much work can be avoided if existing protocols
   or implementations can be readily reused.

18.  Security Considerations

   Security has become a high priority in the design and operation of
   the Internet.  The Internet is vast, and countless organizations and
   individuals own and operate its various components.  A consensus has
   emerged for what might be called a "security placement principle": a
   security mechanism is most effective when it is placed as close as
   possible to, and under the direct control of the owner of the asset
   that it protects.

   A corollary of this principle is that end-to-end security (e.g.,
   confidentiality, authentication, integrity, and access control)
   cannot be ensured with subnetwork security mechanisms.  Not only are
   end-to-end security mechanisms much more closely associated with the
   end-user assets they protect, they are also much more comprehensive.
   For example, end-to-end security mechanisms cover gaps that can
   appear when otherwise good subnetwork mechanisms are concatenated.
   This is an important application of the end-to-end principle [SRC81].

   Several security mechanisms that can be used end-to-end have already
   been deployed in the Internet and are enjoying increasing use.  The
   most important are the Secure Sockets Layer (SSL) [SSL2] [SSL3] and
   TLS [RFC2246] primarily used to protect web commerce, Pretty Good
   Privacy (PGP) [RFC1991] and S/MIME [RFCs-2630-2634], primarily used
   to protect and authenticate email and software distributions, the
   Secure Shell (SSH), used for secure remote access and file transfer,
   and IPsec [RFC2401], a general purpose encryption and authentication
   mechanism that sits just above IP and can be used by any IP
   application.  (IPsec can actually be used either on an end-to-end
   basis or between security gateways that do not include either or both
   end systems.)

   Nonetheless, end-to-end security mechanisms are not used as widely as
   might be desired.  However, the group could not reach consensus on
   whether subnetwork designers should be actively encouraged to
   implement mechanisms to protect user data.

   The clear consensus of the working group held that subnetwork
   security mechanisms, especially when weak or incorrectly implemented
   [BGW01], may actually be counterproductive.  The argument is that
   subnetwork security mechanisms can lull end users into a false sense
   of security, diminish the incentive to deploy effective end-to-end





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   mechanisms, and encourage "risky" uses of the Internet that would not
   be made if users understood the inherent limits of subnetwork
   security mechanisms.

   The other point of view encourages subnetwork security on the
   principle that it is better than the default situation, which all too
   often is no security at all.  Users of especially vulnerable subnets
   (such as consumers who have wireless home networks and/or shared
   media Internet access) often have control over at most one endpoint
   -- usually a client -- and therefore cannot enforce the use of end-
   to-end mechanisms.  However, subnet security can be entirely adequate
   for protecting low-valued assets against the most likely threats.  In
   any event, subnet mechanisms do not preclude the use of end-to-end
   mechanisms, which are typically used to protect highly-valued assets.
   This viewpoint recognizes that many security policies implicitly
   assume that the entire end-to-end path is composed of a series of
   concatenated links that are nominally physically secured.  That is,
   these policies assume that all endpoints of all links are trusted,
   and that access to the physical medium by attackers is difficult.  To
   meet the assumptions of such policies, explicit mechanisms are needed
   for links (especially shared medium links) that lack physical
   protection.  This, for example, is the rationale that underlies Wired
   Equivalent Privacy (WEP) in the IEEE 802.11 [IEEE80211] wireless LAN
   standard, and the Baseline Privacy Interface in the DOCSIS [DOCSIS1]
   [DOCSIS2] data over cable television networks standards.

   We therefore recommend that subnetwork designers who choose to
   implement security mechanisms to protect user data be as candid as
   possible with the details of such security mechanisms and the
   inherent limits of even the most secure mechanisms when implemented
   in a subnetwork rather than on an end-to-end basis.

   In keeping with the "placement principle", a clear consensus exists
   for another subnetwork security role: the protection of the
   subnetwork itself.  Possible threats to subnetwork assets include
   theft of service and denial of service; shared media subnets tend to
   be especially vulnerable to such attacks.  In some cases, mechanisms
   that protect subnet assets can also improve (but cannot ensure) end-
   to-end security.

   One security service can be provided by the subnetwork that will aid
   in the solution of an overall Internet problem: subnetwork security
   should provide a mechanism to authenticate the source of a subnetwork
   frame.  This function is missing in some current protocols, e.g., the
   use of ARP [RFC826] to associate an IPv4 address with a MAC address.
   The IPv6 Neighbor Discovery (ND) [RFC2461] performs a similar
   function.




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   There are well-known security flaws with this address resolution
   mechanism [Wilbur89].  However, the inclusion of subnetwork frame
   source authentication will permit a secure subnetwork address.

   Another potential role for subnetwork security is to protect users
   against traffic analysis, i.e., identifying the communicating parties
   and determining their communication patterns and volumes even when
   their actual contents are protected by strong end-to-end security
   mechanisms.  Lower-layer security can be more effective against
   traffic analysis due to its inherent ability to aggregate the
   communications of multiple parties sharing the same physical
   facilities while obscuring higher-layer protocol information that
   indicates specific end points, such as IP addresses and TCP/UDP port
   numbers.

   However, traffic analysis is a notoriously subtle and difficult
   threat to understand and defeat, far more so than threats to
   confidentiality and integrity.  We therefore urge extreme care in the
   design of subnetwork security mechanisms specifically intended to
   thwart traffic analysis.

   Subnetwork designers must keep in mind that design and implementation
   for security is difficult [Schneier00].  [Schneier95] describes
   protocols and algorithms which are considered well-understood and
   believed to be sound.

   Poor design process, subtle design errors and flawed implementation
   can result in gaping vulnerabilities.  In recent years, a number of
   subnet standards have had problems exposed.  The following are
   examples of mistakes that have been made:

   1.  Use of weak and untested algorithms [Crypto9912] [BGW01].  For a
       variety of reasons, algorithms were chosen which had subtle
       flaws, making them vulnerable to a variety of attacks.

   2.  Use of "security by obscurity" [Schneier00] [Crypto9912].  One
       common mistake is to assume that keeping cryptographic algorithms
       secret makes them more secure.  This is intuitive, but wrong.
       Full public disclosure early in the design process attracts peer
       review by knowledgeable cryptographers.  Exposure of flaws by
       this review far outweighs any imagined benefit from forcing
       attackers to reverse engineer security algorithms.

   3.  Inclusion of trapdoors [Schneier00] [Crypto9912].  Trapdoors are
       flaws surreptitiously left in an algorithm to allow it to be
       broken.  This might be done to recover lost keys or to permit
       surreptitious access by governmental agencies.  Trapdoors can be
       discovered and exploited by malicious attackers.



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   4.  Sending passwords or other identifying information as clear text.
       For many years, analog cellular telephones could be cloned and
       used to steal service.  The cloners merely eavesdropped on the
       registration protocols that exchanged everything in clear text.

   5.  Keys which are common to all systems on a subnet [BGW01].

   6.  Incorrect use of a sound mechanism.  For example [BGW01], one
       subnet standard includes an initialization vector which is poorly
       designed and poorly specified.  A determined attacker can easily
       recover multiple ciphertexts encrypted with the same key stream
       and perform statistical attacks to decipher them.

   7.  Identifying information sent in clear text that can be resolved
       to an individual, identifiable device.  This creates a
       vulnerability to attacks targeted to that device (or its owner).

   8.  Inability to renew and revoke shared secret information.

   9.  Insufficient key length.

   10. Failure to address "man-in-the-middle" attacks, e.g., with mutual
       authentication.

   11. Failure to provide a form of replay detection, e.g., to prevent a
       receiver from accepting packets from an attacker that simply
       resends previously captured network traffic.

   12. Failure to provide integrity mechanisms when providing
       confidentiality schemes [Bel98].

   This list is by no means comprehensive.  Design problems are
   difficult to avoid, but expert review is generally invaluable in
   avoiding problems.

   In addition, well-designed security protocols can be compromised by
   implementation defects.  Examples of such defects include use of
   predictable pseudo-random numbers [RFC1750], vulnerability to buffer
   overflow attacks due to unsafe use of certain I/O system calls
   [WFBA2000], and inadvertent exposure of secret data.

19.  Contributors

   This document represents a consensus of the members of the IETF
   Performance Implications of Link Characteristics (PILC) working
   group.





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   This document would not have been possible without the contributions
   of a great number of people in the Performance Implications of Link
   Characteristics Working Group.  In particular, the following people
   provided major contributions of text, editing, and advice on this
   document: Mark Allman provided the final editing to complete this
   document.  Carsten Bormann provided text on robust header
   compression.  Gorry Fairhurst provided text on broadcast and
   multicast issues, routing,  and many valuable comments on the entire
   document.  Aaron Falk provided text on bandwidth on demand.  Dan
   Grossman provided text on many facets of the document.  Reiner Ludwig
   provided thorough document review and text on TCP vs. Link-Layer
   Retransmission.  Jamshid Mahdavi provided text on TCP performance
   calculations.  Saverio Mascolo provided feedback on the document.
   Gabriel Montenegro provided feedback on the document.  Marie-Jose
   Montpetit provided text on bandwidth on demand.  Joe Touch provided
   text on multicast, broadcast, and routing, and Lloyd Wood provided
   many valuable comments on versions of the document.

20.  Informative References

   References of the form RFCnnnn are Internet Request for Comments
   (RFC) documents available online at www.rfc-editor.org.

   [802.1D]      Information Technology Telecommunications and
                 information exchange between systems Local and
                 metropolitan area networks, Common specifications Media
                 access control (MAC) bridges, IEEE 802.1D, 1998.  ISO
                 15802-3.

   [802.1p]      IEEE, 802.1p, Standard for Local and Metropolitan Area
                 Networks - Supplement to Media Access Control (MAC)
                 Bridges: Traffic Class Expediting and Multicast.

   [AP99]        Allman, M. and V. Paxson, On Estimating End-to-End
                 Network Path Properties, In Proceedings of ACM SIGCOMM
                 99.

   [AR02]        Acar, G. and C. Rosenberg, Weighted Fair Bandwidth-on-
                 Demand (WFBoD) for Geo-Stationary Satellite Networks
                 with On-Board Processing, Computer Networks, 39(1),
                 2002.

   [ATMFTM]      The ATM Forum, "Traffic Management Specification,
                 Version 4.0", April 1996, document af-tm-0056.000.
                 http://www.atmforum.com/






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   [BA02]        Blanton, E. and M. Allman, On Making TCP More Robust to
                 Packet Reordering. ACM Computer Communication Review,
                 32(1), January 2002.

   [Bel98]       Bellovin, S., "Cryptography and the Internet", in
                 Proceedings of CRYPTO '98, August 1998.
                 http://www.research.att.com/~smb/papers/inet-crypto.pdf

   [BGW01]       Borisov, N., Goldberg, I. and D. Wagner, "Intercepting
                 Mobile Communications: The Insecurity of 802.11," In
                 Proceedings of ACM MobiCom, July 2001.

   [BPK98]       Balakrishnan, H., Padmanabhan, V. and R. Katz.  "The
                 Effects of Asymmetry on TCP Performance."  ACM Mobile
                 Networks and Applications (MONET), 1998.

   [BPS99]       Bennet,, J.C.R., Partridge, C. and N. Shectman, "Packet
                 Reordering is Not Pathological Network Behavior",
                 IEEE/ACM Transactions on Networking, Vol. 7, No. 6,
                 December 1999.

   [CGMP]        Farinacci D., Tweedly A. and T. Speakman, "Cisco Group
                 Management Protocol (CGMP)", 1996/1997.
                 ftp://ftpeng.cisco.com/ipmulticast/specs/cgmp.txt

   [Crypto9912]  Schneier, B., "European Cellular Encryption Algorithms"
                 Crypto-Gram, December 15, 1999.
                 http://www.counterpane.com

   [DIX82]       Digital Equipment Corp, Intel Corp, Xerox Corp,
                 Ethernet Local Area Network Specification Version 2.0,
                 November 1982.

   [DOCSIS1]     Data-Over-Cable Service Interface Specifications, Radio
                 Frequency Interface Specification 1.0, SP-RFI-I05-
                 991105, November 1999, Cable Television Laboratories,
                 Inc.

   [DOCSIS2]     Data-Over-Cable Service Interface Specifications, Radio
                 Frequency Interface Specification 1.1, SP-RFIv1.1-I05-
                 000714, July 2000, Cable Television Laboratories, Inc.

   [DOCSIS3]     Lai, W.S., "DOCSIS-Based Cable Networks: Impact of
                 Large Data Packets on Upstream Capacity", 14th ITC
                 Specialists Seminar on Access Networks and Systems,
                 Barcelona, Spain, April 25-27, 2001.





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   [EN301192]    ETSI, European Broadcasting Union, Digital Video
                 Broadcasting (DVB); DVB Specification for Data
                 Broadcasting, European Standard (Telecommunications
                 Series)  EN 301 192 v1.2.1(1999-06).

   [ES00]        Eckhardt, D. and P. Steenkiste, "Effort-limited Fair
                 (ELF) Scheduling for Wireless Networks, Proceedings of
                 IEEE Infocom 2000.

   [FB00]        Firoiu V. and M. Borden, "A Study of Active Queue
                 Management for Congestion Control" to appear in Infocom
                 2000.

   [GM02]        Grieco1, L. and S. Mascolo, "TCP Westwood and Easy RED
                 to Improve Fairness in High-Speed Networks",
                 Proceedings of the 7th International Workshop on
                 Protocols for High-Speed Networks, April 2002.

   [IEEE8023]    IEEE 802.3 CSMA/CD Access Method.
                 http://standards.ieee.org/

   [IEEE80211]   IEEE 802.11 Wireless LAN standard.
                 http://standards.ieee.org/

   [ISO3309]     ISO/IEC 3309:1991(E), "Information Technology -
                 Telecommunications and information exchange between
                 systems - High-level data link control (HDLC)
                 procedures - Frame structure", International
                 Organization For Standardization, Fourth edition 1991-
                 06-01.

   [ISO13818]    ISO/IEC, ISO/IEC 13818-1:2000(E)  Information
                 Technology - Generic coding of moving pictures and
                 associated audio information:  Systems, Second edition,
                 2000-12-01 International Organization for
                 Standardization and International Electrotechnical
                 Commission.

   [ITU-I363]    ITU-T I.363.5 B-ISDN ATM Adaptation Layer Specification
                 Type AAL5, International Standards Organisation (ISO),
                 1996.

   [Jac90]       Jacobson, V., Modified TCP Congestion Avoidance
                 Algorithm.  Email to the end2end-interest mailing list,
                 April 1990.
                 ftp://ftp.ee.lbl.gov/email/vanj.90apr30.txt





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   [KY02]        Khafizov, F. and M. Yavuz, Running TCP Over IS-2000,
                 Proceedings of IEEE ICC, 2002.

   [LK00]        Ludwig, R. and R. H. Katz, "The Eifel Algorithm: Making
                 TCP Robust Against Spurious Retransmissions", ACM
                 Computer Communication Review, Vol. 30, No. 1, January
                 2000.

   [LKJK02]      Ludwig, R., Konrad, A., Joseph, A. D. and R. H. Katz,
                 "Optimizing the End-to-End Performance of Reliable
                 Flows over Wireless Links", Kluwer/ACM Wireless
                 Networks Journal, Vol. 8, Nos. 2/3, pp. 289-299,
                 March-May 2002.

   [LRKOJ99]     Ludwig, R., Rathonyi, B., Konrad, A., Oden, K. and A.
                 Joseph, Multi-Layer Tracing of TCP over a Reliable
                 Wireless Link, pp. 144-154, In Proceedings of ACM
                 SIGMETRICS 99.

   [LS00]        Ludwig, R. and K. Sklower, The Eifel Retransmission
                 Timer, ACM Computer Communication Review, Vol. 30, No.
                 3, July 2000.

   [MAGMA-PROXY] Fenner, B., He, H., Haberman, B. and H. Sandick,
                 "IGMP/MLD-based Multicast Forwarding ("IGMP/MLD
                 Proxying")", Work in Progress.

   [MAGMA-SNOOP] Christensen, M., Kimball, K. and F. Solensky,
                 "Considerations for IGMP and MLD Snooping Switches",
                 Work in Progress.

   [MBB00]       May, M., Bonald, T. and J-C. Bolot, "Analytic
                 Evaluation of RED Performance", INFOCOM 2000.

   [MBDL99]      May, M., Bolot, J., Diot, C. and B. Lyles, "Reasons not
                 to deploy RED", Proc. of 7th. International Workshop on
                 Quality of Service (IWQoS'99), June 1999.

   [MSMO97]      Mathis, M., Semke, J., Mahdavi, J. and T. Ott, "The
                 Macroscopic Behavior of the TCP Congestion Avoidance
                 Algorithm", Computer Communication Review, Vol. 27,
                 number 3, July 1997.

   [MYR95]       Boden, N., Cohen, D., Felderman, R., Kulawik, A.,
                 Seitz, C., et al.  MYRINET: A Gigabit per Second Local
                 Area Network, IEEE-Micro, Vol. 15, No.1, February 1995,
                 pp. 29-36.




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   [PFTK98]      Padhye, J., Firoiu, V., Towsley, D. and J. Kurose,
                 "Modeling TCP Throughput: a Simple Model and its
                 Empirical Validation", UMASS CMPSCI Tech Report TR98-
                 008, Feb. 1998.

   [RED93]       Floyd, S. and V. Jacobson, "Random Early Detection
                 gateways for Congestion Avoidance", IEEE/ACM
                 Transactions in Networking, Vol. 1 No. 4, August 1993.
                 http://www.aciri.org/floyd/papers/red/red.html

   [RF95]        Romanow, A. and S. Floyd, "Dynamics of TCP Traffic over
                 ATM Networks".  IEEE Journal of Selected Areas in
                 Communication, Vol.13 No.  4, May 1995, p. 633-641.

   [RFC791]      Postel, J., "Internet Protocol", STD 5, RFC 791,
                 September 1981.

   [RFC793]      Postel, J., "Transmission Control Protocol", STD 7, RFC
                 793, September 1981.

   [RFC768]      Postel, J., "User Datagram Protocol", STD 6, RFC 768,
                 August 1980.

   [RFC826]      Plummer, D.C., "Ethernet Address Resolution Protocol:
                 Or converting network protocol addresses to 48-bit
                 Ethernet address for transmission on Ethernet
                 hardware", STD 37, RFC 826, November 1982.

   [RFC1071]     Braden, R., Borman, D. and C. Partridge, "Computing the
                 Internet checksum", RFC 1071, September 1988.

   [RFC1112]     Deering, S., "Host Extensions for IP Multicasting", STD
                 5, RFC 1112, August 1989.

   [RFC1144]     Jacobson, V., "Compressing TCP/IP Headers for Low-Speed
                 Serial Links", RFC 1144, February 1990.

   [RFC1191]     Mogul, J. and S. Deering, "Path MTU Discovery", RFC
                 1191, November 1990.

   [RFC1332]     McGregor, C., "The PPP Internet Protocol Control
                 Protocol (IPCP)", RFC 1332, May 1992.

   [RFC1435]     Knowles, S., "IESG Advice from Experience with Path MTU
                 Discovery", RFC 1435, March 1993.






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   [RFC1633]     Braden, R., Clark, D. and S. Shenker, "Integrated
                 Services in the Internet Architecture: an Overview",
                 RFC 1633, June 1994.

   [RFC1661]     Simpson, W., "The Point-to-Point Protocol (PPP)", STD
                 51, RFC 1661, July 1994.

   [RFC1662]     Simpson, W., Ed., "PPP in HDLC-like Framing", STD 51,
                 RFC 1662, July 1994.

   [RFC1750]     Eastlake 3rd, D., Crocker, S. and J. Schiller,
                 "Randomness Recommendations for Security", RFC 1750,
                 December 1994.

   [RFC1812]     Baker, F., Ed., "Requirements for IP Version 4
                 Routers", RFC 1812, June 1995.

   [RFC1939]     Myers, J. and M. Rose, "Post Office Protocol - Version
                 3", STD 53, RFC 1939, May 1996.

   [RFC1981]     McCann, J., Deering, S. and J. Mogul, "Path MTU
                 Discovery for IP version 6", RFC 1981, August 1996.

   [RFC1991]     Atkins, D., Stallings, W. and P. Zimmermann, "PGP
                 Message Exchange Formats", RFC 1991, August 1996.

   [RFC2018]     Mathis, M., Mahdavi, J., Floyd, S. and A. Romanow, "TCP
                 Selective Acknowledgement Options", RFC 2018, October
                 1996.

   [RFC2131]     Droms, R., "Dynamic Host Configuration Protocol", RFC
                 2131, March 1997.

   [RFC2205]     Braden, R., Ed., Zhang, L., Berson, S., Herzog, S. and
                 S. Jamin, "Resource ReSerVation Protocol (RSVP) --
                 Version 1 Functional Specification", RFC 2205,
                 September 1997.

   [RFC2208]     Mankin, A., Baker, F., Braden, B., Bradner, S., O`Dell,
                 M., Romanow, A., Weinrib, A. and L. Zhang, "Resource
                 ReSerVation Protocol (RSVP) -- Version 1 Applicability
                 Statement Some Guidelines on Deployment", RFC 2208,
                 September 1997.

   [RFC2210]     Wroclawski, J., "The Use of RSVP with IETF Integrated
                 Services", RFC 2210, September 1997.





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   [RFC2211]     Wroclawski, J., "Specification of the Controlled-Load
                 Network Element Service", RFC 2211, September 1997.

   [RFC2212]     Shenker, S., Partridge, C. and R. Guerin,
                 "Specification of Guaranteed Quality of Service", RFC
                 2212, September 1997.

   [RFC2246]     Dierks, T. and C. Allen, "The TLS Protocol Version
                 1.0", RFC 2246, January 1999.

   [RFC2309]     Braden, B., Clark, D., Crowcroft, J., Davie, B.,
                 Deering, S., Estrin, D., Floyd, S., Jacobson, V.,
                 Minshall, G., Partridge, C., Peterson, L.,
                 Ramakrishnan, K., Shenker, S., Wroclawski, J. and L.
                 Zhang, "Recommendations on Queue Management and
                 Congestion Avoidance in the Internet", RFC 2309, April
                 1998.

   [RFC2322]     van den Hout, K., Koopal, A. and R. van Mook,
                 "Management of IP numbers by peg-dhcp", RFC 2322, 1
                 April 1998.

   [RFC2328]     Moy, J., "OSPF Version 2", STD 54, RFC 2328, April
                 1998.

   [RFC2332]     Luciani, J., Katz, D., Piscitello, D., Cole, B. and N.
                 Doraswamy, "NBMA Next Hop Resolution Protocol (NHRP)",
                 RFC 2332, April 1998.

   [RFC2364]     Gross, G., Kaycee, M., Li, A., Malis, A. and J.
                 Stephens, "PPP Over AAL5", RFC 2364, July 1998.

   [RFC2394]     Pereira, R., "IP Payload Compression Using DEFLATE",
                 RFC 2394, December 1998.

   [RFC2395]     Friend, R. and R. Monsour, "IP Payload Compression
                 Using LZS", RFC 2395, December 1998.

   [RFC2401]     Kent, S. and R. Atkinson, "Security Architecture for
                 the Internet Protocol", RFC 2401, November 1998.

   [RFC2406]     Kent, S. and R. Atkinson, "IP Encapsulating Security
                 Payload (ESP)", RFC 2406, November 1998.

   [RFC2440]     Callas, J., Donnerhacke, L., Finney, H. and R. Thayer,
                 "OpenPGP Message Format", RFC 2440, November 1998.





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   [RFC2460]     Deering, S. and R. Hinden, "Internet Protocol, Version
                 6 (IPv6) Specification", RFC 2460, December 1998.

   [RFC2461]     Narten, T., Nordmark, E. and W. Simpson, "Neighbor
                 Discovery for IP Version 6 (IPv6)", RFC 2461, December
                 1998.

   [RFC2474]     Nichols, K., Blake, S., Baker, F. and D. Black,
                 "Definition of the Differentiated Services Field (DS
                 Field) in the IPv4 and IPv6 Headers", RFC 2474,
                 December 1998.

   [RFC2475]     Blake, S., Black, D., Carlson, M., Davies, E., Wang, Z.
                 and W. Weiss, "An Architecture for Differentiated
                 Services", RFC 2475, December 1998.

   [RFC2507]     Degermark, M., Nordgren, B. and S. Pink, "IP Header
                 Compression", RFC 2507, February 1999.

   [RFC2508]     Casner, S. and V. Jacobson, "Compressing IP/UDP/RTP
                 Headers for Low-Speed Serial Links", RFC 2508, February
                 1999.

   [RFC2581]     Allman, M., Paxson, V. and W. Stevens, "TCP Congestion
                 Control", RFC 2581, April 1999.

   [RFC2582]     Floyd, S. and T. Henderson, "The NewReno Modification
                 to TCP's Fast Recovery Algorithm", RFC 2582, April
                 1999.

   [RFC2597]     Heinanen, J., Baker, F., Weiss, W. and J. Wroclawski,
                 "Assured Forwarding PHB Group", RFC 2597, June 1999.

   [RFC2616]     Fielding, R., Gettys, J., Mogul, J., Frystyk, H.,
                 Masinter, L., Leach, P. and T. Berners-Lee, "Hypertext
                 Transfer Protocol -- HTTP/1.1", RFC 2616, June 1999.

   [RFC2630]     Housley, R., "Cryptographic Message Syntax", RFC 2630,
                 June 1999.

   [RFC2631]     Rescorla, E., "Diffie-Hellman Key Agreement Method",
                 RFC 2631, June 1999.

   [RFC2632]     Ramsdell, B., Ed., "S/MIME Version 3 Certificate
                 Handling", RFC 2632, June 1999.






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   [RFC2633]     Ramsdell, B., "S/MIME Version 3 Message Specification",
                 RFC 2633, June 1999.

   [RFC2634]     Hoffman, P., "Enhanced Security Services for S/MIME",
                 RFC 2634, June 1999.

   [RFC2684]     Grossman, D. and J. Heinanen, "Multiprotocol
                 Encapsulation over ATM Adaptation Layer 5", RFC 2684,
                 September 1999.

   [RFC2686]     Bormann, C., "The Multi-Class Extension to Multi-Link
                 PPP", RFC 2686, September 1999.

   [RFC2687]     Bormann, C., "PPP in a Real-time Oriented HDLC-like
                 Framing", RFC 2687, September 1999.

   [RFC2689]     Bormann, C., "Providing Integrated Services over Low-
                 bitrate Links", RFC 2689, September 1999.

   [RFC2710]     Deering, S., Fenner, W. and B. Haberman, "Multicast
                 Listener Discovery (MLD) for IPv6", RFC 2710, October
                 1999.

   [RFC2784]     Farinacci, D., Li, T., Hanks, S., Meyer, D. and P.
                 Traina, "Generic Routing Encapsulation (GRE)", RFC
                 2784, March 2000.

   [RFC2865]     Rigney, C., Willens, S., Rubens, A. and W. Simpson,
                 "Remote Authentication Dial In User Service (RADIUS)",
                 RFC 2865, June 2000.

   [RFC2914]     Floyd, S., "Congestion Control Principles", BCP 41, RFC
                 2914, September 2000.

   [RFC2923]     Lahey, K., "TCP Problems with Path MTU Discovery", RFC
                 2923, September 2000.

   [RFC2988]     Paxson, V. and M. Allman, "Computing TCP's
                 Retransmission Timer", RFC 2988, November 2000.

   [RFC2990]     Huston, G., "Next Steps for the IP QoS Architecture",
                 RFC 2990, November 2000.

   [RFC3048]     Whetten, B., Vicisano, L., Kermode, R., Handley, M.,
                 Floyd, S. and M. Luby, "Reliable Multicast Transport
                 Building Blocks for One-to-Many Bulk-Data Transfer",
                 RFC 3048, January 2001.




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   [RFC3095]     Bormann, C., Ed., Burmeister, C., Degermark, M.,
                 Fukushima, H., Hannu, H., Jonsson, L-E., Hakenberg, R.,
                 Koren, T., Le, K., Liu, Z., Martensson, A., Miyazaki,
                 A., Svanbro, K., Wiebke, T., Yoshimura, T. and H.
                 Zheng, "RObust Header Compression (ROHC):  Framework
                 and four profiles: RTP, UDP, ESP, and uncompressed",
                 RFC 3095, July 2001.

   [RFC3096]     Degermark, M., Ed., "Requirements for robust IP/UDP/RTP
                 header compression", RFC 3096, July 2001.

   [RFC3150]     Dawkins, S., Montenegro, G., Kojo, M. and V. Magret,
                 "End-to-end Performance Implications of Slow Links",
                 BCP 48, RFC 3150, July 2001.

   [RFC3155]     Dawkins, S., Montenegro, G., Kojo, M., Magret, V. and
                 N. Vaidya, "End-to-end Performance Implications of
                 Links with Errors", BCP 50, RFC 3155, August 2001.

   [RFC3168]     Ramakrishnan, K., Floyd, S. and D. Black, "The Addition
                 of Explicit Congestion Notification (ECN) to IP", RFC
                 3168, September 2001.

   [RFC3173]     Shacham, A., Monsour, B., Pereira, R. and M. Thomas,
                 "IP Payload Compression Protocol (IPComp)", RFC 3173,
                 September 2001.

   [RFC3246]     Davie, B., Charny, A., Bennet, J.C.R., Benson, K., Le
                 Boudec, J.Y., Courtney, W., Davari, S., Firoiu, V. and
                 D. Stiliadis, "An Expedited Forwarding PHB (Per-Hop
                 Behavior)", RFC 3246, March 2002.

   [RFC3248]     Armitage, G., Carpenter, B., Casati, A., Crowcroft, J.,
                 Halpern, J., Kumar, B. and J. Schnizlein, "A Delay
                 Bound alternative revision of RFC 2598", RFC 3248,
                 March 2002.

   [RFC3344]     Perkins, C., Ed., "IP Mobility Support for IPv4", RFC
                 3344, August 2002.

   [RFC3366]     Fairhurst, G. and L. Wood, "Advice to link designers on
                 link Automatic Repeat reQuest (ARQ)", BCP 62, RFC 3366,
                 August 2002.








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   [RFC3376]     Cain, B., Deering, S., Kouvelas, I., Fenner, B. and A.
                 Thyagarajan, "Internet Group Management Protocol,
                 Version 3", RFC 3376, October 2002.

   [RFC3449]     Balakrishnan, H., Padmanabhan, V., Fairhurst, G. and M.
                 Sooriyabandara, "TCP Performance Implications of
                 Network Path Asymmetry", BCP 69, RFC 3449, December
                 2002.

   [RFC3450]     Luby, M., Gemmell, J., Vicisano, L., Rizzo, L. and J.
                 Crowcroft, "Asynchronous Layered Coding (ALC) Protocol
                 Instantiation", RFC 3450, December 2002.

   [RFC3451]     Luby, M., Gemmell, J., Vicisano, L., Rizzo, L.,
                 Handley, M. and J. Crowcroft, "Layered Coding Transport
                 (LCT) Building Block", RFC 3451, December 2002.

   [RFC3452]     Luby, M., Vicisano, L., Gemmell, J., Rizzo, L.,
                 Handley, M. and J. Crowcroft, "Forward Error Correction
                 (FEC) Building Block", RFC 3452, December 2002.

   [RFC3453]     Luby, M., Vicisano, L., Gemmell, J., Rizzo, L.,
                 Handley, M. and J. Crowcroft, "The Use of Forward Error
                 Correction (FEC) in Reliable Multicast", RFC 3453,
                 December 2002.

   [RFC3488]     Wu, I. and T. Eckert, "Cisco Systems Router-port Group
                 Management Protocol (RGMP)", RFC 3488, February 2003.

   [RFC3501]     Crispin, M., "INTERNET MESSAGE ACCESS PROTOCOL -
                 VERSION 4rev1", RFC 3501, March 2003.

   [RFC3828]     Larzon, L-A., Degermark, M., Pink, S., Jonsson, L-E.,
                 Ed. and G. Fairhurst, Ed., "The User Datagram Protocol
                 (UDP)-Lite Protocol", RFC 3828, June 2004.

   [Schneier95]  Schneier, B., Applied Cryptography: Protocols,
                 Algorithms and Source Code in C (John Wiley and Sons,
                 October 1995).

   [Schneier00]  Schneier, B., Secrets and Lies: Digital Security in a
                 Networked World (John Wiley and Sons, August 2000).

   [SP2000]      Stone, J. and C. Partridge, "When the CRC and TCP
                 Checksum Disagree", ACM SIGCOMM, September 2000.
                 http://www.acm.org/sigcomm/sigcomm2000/conf/
                 paper/sigcomm2000-9-1.pdf




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   [SRC81]       Saltzer, J., Reed D. and D. Clark, "End-to-End
                 Arguments in System Design".  Second International
                 Conference on Distributed Computing Systems (April,
                 1981) pages 509-512. Published with minor changes in
                 ACM Transactions in Computer Systems 2, 4, November,
                 1984, pages 277-288. Reprinted in Craig Partridge,
                 editor Innovations in internetworking. Artech House,
                 Norwood, MA, 1988, pages 195-206. ISBN 0-89006-337-0.

   [SSL2]        Hickman, K., "The SSL Protocol", Netscape
                 Communications Corp., Feb 9, 1995.

   [SSL3]        Frier, A., Karlton, P. and P. Kocher, "The SSL 3.0
                 Protocol", Netscape Communications Corp., Nov 18, 1996.

   [TCPF98]      Lin, D. and H.T. Kung, "TCP Fast Recovery Strategies:
                 Analysis and Improvements", IEEE Infocom, March 1998.
                 http://www.eecs.harvard.edu/networking/papers/infocom-
                 tcp-final-198.pdf

   [WFBA2000]    Wagner, D., Foster, J., Brewer, E. and A. Aiken, "A
                 First Step Toward Automated Detection of Buffer Overrun
                 Vulnerabilities", Proceedings of NDSS2000.
                 http://www.isoc.org/isoc/conferences/ndss/
                 2000/proceedings/039.pdf

   [Wilbur89]    Wilbur, Steve R., Jon Crowcroft, and Yuko Murayama.
                 "MAC layer Security Measures in Local Area Networks",
                 Local Area Network Security, Workshop LANSEC '89
                 Proceedings, Springer-Verlag, April 1989, pp. 53-64.





















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21. Contributors' Addresses

   Aaron Falk
   USC/Information Sciences Institute
   4676 Admiralty Way
   Marina Del Rey, CA 90292

   Phone: 310-448-9327
   EMail: falk@isi.edu


   Saverio Mascolo
   Dipartimento di Elettrotecnica ed Elettronica,
   Politecnico di Bari Via Orabona 4, 70125 Bari, Italy

   Phone: +39 080 596 3621
   EMail: mascolo@poliba.it
   URL: http://www-dee.poliba.it/dee-web/Personale/mascolo.html


   Marie-Jose Montpetit
   MJMontpetit.com

   EMail: marie@mjmontpetit.com



























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22.  Authors' Addresses

   Phil Karn, Editor
   Qualcomm 5775 Morehouse Drive
   San Diego CA 92121

   Phone: 858 587 1121
   EMail: karn@qualcomm.com


   Carsten Bormann
   Universitaet Bremen TZI
   Postfach 330440
   D-28334 Bremen, Germany

   Phone: +49 421 218 7024
   Fax:   +49 421 218 7000
   EMail: cabo@tzi.org


   Godred (Gorry) Fairhurst
   Department of Engineering, University of Aberdeen,
   Aberdeen, AB24 3UE, United Kingdom

   EMail: gorry@erg.abdn.ac.uk
   URL: http://www.erg.abdn.ac.uk/users/gorry


   Dan Grossman
   Motorola, Inc.
   111 Locke Drive
   Marlboro, MA 01752

   EMail: Dan.Grossman@motorola.com


   Reiner Ludwig
   Ericsson Research
   Ericsson Allee
   1 52134 Herzogenrath, Germany

   Phone: +49 2407 575 719
   EMail: Reiner.Ludwig@ericsson.com








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   Jamshid Mahdavi
   Novell, Inc.

   EMail: jmahdavi@earthlink.net


   Gabriel Montenegro
   Sun Microsystems Laboratories, Europe
   180, Avenue de l'Europe
   38334 Saint Ismier CEDEX
   France

   EMail: gab@sun.com


   Joe Touch
   USC/Information Sciences Institute
   4676 Admiralty Way
   Marina del Rey CA 90292

   Phone: 310 448 9151
   EMail: touch@isi.edu
   URL: http://www.isi.edu/touch


   Lloyd Wood
   Cisco Systems
   9 New Square Park, Bedfont Lakes
   Feltham TW14 8HA
   United Kingdom

   Phone: +44 (0)20 8824 4236
   EMail: lwood@cisco.com
   URL: http://www.ee.surrey.ac.uk/Personal/L.Wood/

















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23.  Full Copyright Statement

   Copyright (C) The Internet Society (2004).  This document is subject
   to the rights, licenses and restrictions contained in BCP 78, and
   except as set forth therein, the authors retain all their rights.

   This document and the information contained herein are provided on an
   "AS IS" basis and THE CONTRIBUTOR, THE ORGANIZATION HE/SHE
   REPRESENTS OR IS SPONSORED BY (IF ANY), THE INTERNET SOCIETY AND THE
   INTERNET ENGINEERING TASK FORCE DISCLAIM ALL WARRANTIES, EXPRESS OR
   IMPLIED, INCLUDING BUT NOT LIMITED TO ANY WARRANTY THAT THE USE OF
   THE INFORMATION HEREIN WILL NOT INFRINGE ANY RIGHTS OR ANY IMPLIED
   WARRANTIES OF MERCHANTABILITY OR FITNESS FOR A PARTICULAR PURPOSE.

Intellectual Property

   The IETF takes no position regarding the validity or scope of any
   Intellectual Property Rights or other rights that might be claimed
   to pertain to the implementation or use of the technology
   described in this document or the extent to which any license
   under such rights might or might not be available; nor does it
   represent that it has made any independent effort to identify any
   such rights.  Information on the procedures with respect to
   rights in RFC documents can be found in BCP 78 and BCP 79.

   Copies of IPR disclosures made to the IETF Secretariat and any
   assurances of licenses to be made available, or the result of an
   attempt made to obtain a general license or permission for the use
   of such proprietary rights by implementers or users of this
   specification can be obtained from the IETF on-line IPR repository
   at http://www.ietf.org/ipr.

   The IETF invites any interested party to bring to its attention
   any copyrights, patents or patent applications, or other
   proprietary rights that may cover technology that may be required
   to implement this standard.  Please address the information to the
   IETF at ietf-ipr@ietf.org.

Acknowledgement

   Funding for the RFC Editor function is currently provided by the
   Internet Society.









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