Internet Engineering Task Force                                Phil Karn
INTERNET DRAFT                                                Aaron Falk
                                                               Joe Touch
                                                    Marie-Jose Montpetit
                                                         Jamshid Mahdavi
                                                      Gabriel Montenegro

File: draft-ietf-pilc-link-design-01.txt                   October, 1999
                                                 Expires:    April, 2000

                Advice for Internet Subnetwork Designers

Status of this Memo

   This document is an Internet-Draft and is in full conformance with
   all provisions of Section 10 of RFC2026.

   Internet-Drafts are working documents of the Internet Engineering
   Task Force (IETF), its areas, and its working groups.  Note that
   other groups may also distribute working documents as Internet-

   Internet-Drafts are draft documents valid for a maximum of six months
   and may be updated, replaced, or obsoleted by other documents at any
   time.  It is inappropriate to use Internet- Drafts as reference
   material or to cite them other than as "work in progress."

   The list of current Internet-Drafts can be accessed at

   The list of Internet-Draft Shadow Directories can be accessed at


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

   This document represents an evolving consensus of the members of the
   IETF Performance Implications of Link Characteristics (PILC) working

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Introduction and Overview

   The Internet Protocol [RFC791] is the core protocol of the world-wide
   Internet that defines a simple "connectionless" packet-switched
   network.  The success of the Internet is largely attributed to the
   simplicity of IP, the "end-to-end principle" 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.

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

   With the explosive growth of the Internet, IP is 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

   Optimizing a subnetwork for IP involves three complementary

   1. Providing functionality sufficient to carry IP.

   2. Eliminating unnecessary functions that increase cost or

   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. I.e., subnetwork designers make many more errors of
   commission than errors of omission.  But certain enhanced Internet
   features, such as multicasting and quality-of-service, rely on
   support from the underlying subnetworks beyond that necessary to
   carry "traditional" unicast, best-effort IP.

   A major consideration in the efficient design of any layered
   communication network are the appropriate layer(s) in which to
   implement a given feature. This issue was first addressed in the
   seminal paper "End-to-End Arguments in System Design" [SRC81]. This
   paper argued that many -- if not most -- network functions are best
   implemented on an end-to-end basis, i.e., at the higher protocol
   layers.  Duplicating these functions at the lower levels is usually
   redundant, and can even be harmful. However, certain low level
   functions can sometimes be justified as a performance enhancement.

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   An example would be link layer retransmission on an unusually lossy
   channel, e.g., mobile radio.

   The architecture of the Internet was heavily influenced by the end-
   to-end principle, and in our view it was crucial to the Internet's

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

Maximum Transmission Units (MTUs) and IP Fragmentation

   IP packets (datagrams) vary in size from 20 bytes (the size of the IP
   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 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 lengths of the
   packets they carry.  One example is Ethernet with the widely used DIX
   (not IEEE 802.3) header, which lacks a length field to indicate the
   true data length when the packet is padded to the 60 byte minimum.
   This is not a problem for uncompressed IP because it carries its own
   length field.

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

   In IP version 4 (current IP), 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, fragmentation can occur only at the sending host; it
   cannot occur in a router.

   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
   reducing its packet sizes accordingly. This procedure is optional in
   IPv4 but mandatory in IPv6 where there is no router fragmentation.

   The Path MTU Discovery procedure (and the deletion of router
   fragmentation in IPv6) reflects a consensus of the Internet technical
   community that IP fragmentation is best avoided. This requires that

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   subnetworks support MTUs that are "reasonably" large. The smallest
   MTU that IPv4 can use is 28 bytes, but this is clearly unreasonable;
   because each IP header is 20 bytes, only 8 bytes per packet would be
   available to carry transport headers and application data.

   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.

   This leaves the question of what is a "reasonable" MTU.  Ethernet (10
   and 100 Mb/s) has a MTU of 1500 bytes, and because of its ubiquity
   few Internet paths have MTUs larger than this value.  This severely
   limits the utility of larger MTUs provided by other subnetworks. But
   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.

  Choosing the MTU in Slow Networks [Stevens94, RFC1144]

   In slow networks, the time required to transmit the largest possible
   packet may be considerable.  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,
   queueing 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, and these are in turn directly influenced by the
   maximum transmission unit (MTU). 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 a 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 IP fragmentation (if Path MTU discovery is not in use).

   One way to limit delay for interactive traffic without imposing a
   small MTU is to preempt (abort) 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.

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   Another way 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. ATM (asynchronous transfer mode) is an 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%) additional overhead due to the addition of 5-byte
   frame headers to each 48-byte data frame.

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

Framing on Connection-Oriented Subnetworks

   IP needs a way to 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;


   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] is
   applicable to bit synchronous, octet synchronous and octet
   asynchronous links (i.e., examples 1-3 above). 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.

   In particular, IP packet boundaries may, but 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 (SDU), and to rely on the subnetwork's existing
   ability to delimit SDUs to also delimit IP packets.  A good example

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   is Ethernet. But some subnetworks have SDUs 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 frame size of 48 bytes plus a 5-byte 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 variable sized, they may not necessarily fit
   into an integer multiple of fixed-sized SDUs. An "adaptation layer"
   is needed to convert IP packets into SDUs while marking the boundary
   between each IP packet in some manner.

   There are two traditional approaches to the problem. The first is to
   encode each IP packet into one or more SDUs, with no SDU containing
   pieces of more than one IP packet, and padding out the last SDU of
   the packet as needed.  Bits in a control header added to each SDU
   indicate where it 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 to indicate whether the SDU is the first and/or the
   last in the IP packet. Or only the last SDU of the packet could be
   marked, as this would implicitly mark the next SDU as the first in a
   new IP packet. The AAL5 (ATM Adaption 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.

   The second approach is to insert a special flag sequence into the
   data stream between each IP packet, and to pack the resulting data
   stream into SDUs without regard to SDU boundaries. The flag sequence
   can also pad unused space at the end of an SDU. 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

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

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, the Transmission Control Protocol, running
   atop IP on an end-to-end basis.

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   Connection-oriented subnetworks can be (and are) widely used to carry
   IP, but often with considerable complexity.  Subnetworks with a few
   nodes can simply open a permanent connection between each pair of
   nodes, as is frequently done with ATM. But the number of connections
   is equal to the square of the number of nodes, so this is clearly
   impractical for large subnetworks. A "shim" layer between IP and the
   subnetwork is therefore required to manage connections in the latter.

   These shim layers 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 its call management system
   (especially if was designed for voice traffic 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.

   The ideal subnetwork for IP is connectionless. Connection-oriented
   networks that dedicate minimal resources to each connection (e.g.,
   ATM) are a distant second, and connection-oriented networks that
   dedicate a fixed amount of bandwidth to each connection (e.g., the
   PSTN, including ISDN) are the least efficient. If such subnetworks
   must be used to carry IP, their call-processing systems should be
   capable of rapid call set-up and tear-down.

Bandwidth on Demand (BoD) Subnets (Aaron Falk)

   Wireless networks, including both satellite and terrestrial, may use
   Bandwidth on Demand (BoD). Bandwidth on demand, which is implemented
   at the link layer by Demand Assignment Multiple Access (DAMA) in TDMA
   systems, is currently one of the proposed mechanism to efficiently
   share limited spectrum resources amongst a large number of users.

   The design parameters for BoD are similar to those in connection
   oriented subnetworks, however the implementations may be very
   different. In BoD, the user typically requests access to the shared
   channel for some duration. Access may be allocated in terms of a
   period of time at a specific rate, a certain number of packets, or
   until the user chooses to release the channel. Access may be
   coordinated through a central management entity or through using a
   distributed algorithm amongst the users. The resource shared may be a
   terrestrial wireless hop, a satellite uplink, or an end-to-end
   satellite channel.

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   Long delay BoD subnets pose problems similar to the Connection
   Oriented networks in terms of anticipating traffic arrivals. 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 done by
   piggybacking new requests on user data).

   There are a wide variety of BoD protocols available and there has
   been relatively little comprehensive research on the interactions
   between the BoD mechanisms and Internet protocol performance. A
   tradeoff exists balancing the time a user can be allowed to hold a
   channel to drain port buffers with the additional imposed latency on
   other users who are forced to wait to get access to the channel. It
   is desirable to design mechanisms that constrain the BoD imposed
   latency variation. This will be helpful in preventing spurious
   timeouts from TCP.

Reliability and Error Control

   In the Internet architecture, the ultimate responsibility for error
   recovery is at the end points. 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) must recover from
   these errors on an end-to-end basis.  Error recovery in the
   subnetwork is therefore justified 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 in fact a stated goal
   within the IETF community) 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

 How TCP Works

   One of TCP's functions is end-host based congestion control for the

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   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 (see
   [RED93]), but it is still widely practiced.

   TCP uses sequence numbering and acknowledgements (ACKs) on an end-to-
   end basis to provide reliable, sequenced, once-only delivery.  TCP
   ACKs are cumulative, i.e., each one implicitly ACKs every segment
   received so far.  If a packet is lost, the cumulative ACK will cease
   to advance.

   Since the most common cause of packet loss is congestion, TCP treats
   packet loss as a network congestion indicator. This happens
   automatically, and the subnetwork need not know anything about IP or
   TCP. It simply drops packets whenever it must, though RED shows that
   some packet-dropping strategies are more fair than others.

   TCP recovers from packet losses in two different ways. The most
   important is by a 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 will be idle during the timeout interval and
   restarts with a congestion window of 1 following the timeout. To
   allow faster recovery from the occasional lost packet in a bulk
   transfer, an alternate scheme known as "fast recovery" was introduced

   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, but they will not actually ACK any 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 it can
   retransmit it without waiting for a timeout.  Dupacks effectively
   constitute a negative acknowledgement (NAK) for the packet whose

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   sequence number is equal to the acknowledgement field in the incoming
   TCP packet.  TCP currently waits until a certain number of dupacks
   (currently 3) are seen prior to assuming a loss has occurred; this
   helps avoid an unnecessary retransmission in the face of out-of-
   sequence delivery.

   A new technique called "Explicit Congestion Notification" (ECN)
   allows routers to directly signal congestion to hosts without
   dropping packets.  This is done by setting a bit in the IP header.
   Since this is currently an optional behavior (and, longer term, there
   will always be the possibility of congestion in portions of the
   network which don't support ECN), 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 is the end-system
   congestion control algorithm used by TCP.  This algorithm maintains a
   congestion window (cwnd), which controls the amount of data which TCP
   may have in flight at any given point in time.  Reducing cwnd reduces
   the overall bandwidth obtained by the connection; similarly, raising
   cwnd increases the performance, up to the limit of the available

   TCP probes for available network bandwidth by setting cwnd at one
   packet and then increasing it 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 signalled by other
   mechanisms), cwnd is set back to one and the slow start process is
   repeated until cwnd reaches one half of its previous setting before
   the loss. 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 referred to as an "Additive Increase, Multiplicative
   Decrease" (AIMD) algorithm.  The steep decrease in response to
   congestion provides for network stability; the AIMD algorithm also
   provides for fairness between long running TCP connections sharing
   the same path.

 TCP Performance Characteristics


   In this section, we present the current "state-of-the-art"
   understanding of TCP performance.  This analysis attempts to
   characterize the performance of TCP connections over links of varying

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   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 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 be different than actual performance in the
   Internet.  That being said, we have done the best we can to provide
   information which will help designers get an accurate picture of the
   capabilities and limitations of TCP under various conditions.

  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 is given by Padhye, [PFTK98].  This formula is:

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

   In this formula, the variables are as follows:
           MSS  is the 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)

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

   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,

                     MSS   1
           BW = 0.93 --- -------
                     RTT sqrt(p)

  Assumptions of these formulae

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

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

   The RTT used is the average RTT including queuing delays.

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

   The formulae assume long running TCP connections.  For connections
   which 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

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

  Analysis of Link Layer Effects on TCP Performance

   Link layer designers who are interested in understanding the
   performance of TCP over these links can use these formulae to figure
   this out.  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 layer were used in the Internet, on a path which
   otherwise had a round trip of of 80 msec, you could compute an upper
   bound on the performance as follows:

   For MSS, use 1000 bytes (remove the 40 bytes for TCP/IP headers,
   which do not contribute to performance).

   For RTT, use 120 msec (80 msec for the Internet part, plus 20 msec

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   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

   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" which
   must be accepted.

   2.  Note that the equations for TCP performance are all expressed in
   terms of packet loss.  Many link-layer designers think in terms of
   bit-error rate.  *If* there were a uniform random distribution of
   errors, then the probability of a packet being corrupted would be:

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

   (Here we assume MSS is represented in bytes).  If the inequality

   BER * MSS * 8 << 1

   holds, p can be approximated by:

   p = BER * MSS * 8

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

   Note that links with Forward Error Correction (FEC) generally have
   very non-uniform bit error distributions.  The distribution is a
   strong function of the types and combinations of FEC algorithms used.
   In such cases these equations cannot be used to apply BER to the
   performance equations above.  If the distribution of error
   distributions 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 will need to be used to determine the actual packet

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   loss rate.

   3.  Note that the packet size plays an important role.  Larger packet
   sizes will allow for improved performance at the same *packet loss*
   rate.  Assuming constant, uniform bit-errors (instead of packet
   errors), and assuming that the BER is small enough for the
   approximation [p=BER*MSS*8] to apply, a simple derivation will show
   that larger packet sizes still result in increased TCP performance.
   For this reason (and others) it is advisable to support larger packet
   sizes where possible.

   To derive this, simply plug in p = BER*MSS*8 into the simple formula
   for performance.  The result is p = O(sqrt(MSS)), providing larger
   performance for larger packet sizes.

   If the approximation p = BER*MSS*8 breaks down, and in particular if
   the BER is high enough that BER*MSS approaches (or exceeds) 1, the
   packet loss rate p will tend to 100%, resulting in zero throughput.

   4.  We have chosen a specific RTT which might occur on a wide-area
   Internet path within the USA.  In the Internet, it is important to
   recognize that RTT varies considerably.

   For example, in a wired LAN environment, RTTs are typically less than
   10 msec.  International connections (between hosts in different
   countries) may have RTTs of 200 msec or more.  Modems and other low-
   capacity links can add considerable delay to the overall RTTs
   experienced by the end hosts due to their long packet transmission

   Links running over geostationary repeater satellites have one-way
   times of around 250ms (125ms up to the satellite, 125ms down) so the
   RTT of an end-to-end TCP connection that includes such a link can be
   expected to be greater than 250ms.

   Heavily congested links may have queues which back up, increasing
   RTTs.  Finally, VPNs and other forms of encryption and tunneling can
   add significant end-to-end delay to network connections.

   Increased delay decreases the overall performance of TCP at a given
   loss rate.  A good rule of thumb is to recognize that you can't do
   anything about the laws of physics, so you can't change the
   propagation delay.  Many link layer designers are likely to face the
   following tradeoff: using additional delay to reduce the probability
   of packet loss (through FEC, ARQ, or other methods).  Increasing the
   delay somewhat in order to decrease packet loss is probably a
   worthwhile investment, either up to doubling, or in the case of very
   low delay pipes, adding 10-20 msec won't have much effect on a

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   typical Internet path.

Quality of Service, Fairness vs Performance, Congestion signalling

   [subnet hooks for QOS bits]

Delay Characteristics

   [self clocking TCP, (re)transmission shaping]

Bandwidth Asymmetries

   Some subnetworks may provide asymmetric bandwidth 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 bandwidth of the subnetwork carrying
   the data to the bandwidth of the subnetwork carrying the
   acknowledgments is too large, the slow return of 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 avoid large differences in the
   incoming and outgoing bandwidth.

   One way to cope with asymmetric subnetworks is to increase the size
   of the data segments as much as possible.  This allows more data to
   be sent per ACK, and therefore mitigates the slow flow of ACKs.
   Using the delayed acknowledgment mechanism {Bra89], which reduces the
   number of ACKs transmitted by the receiver by roughly half, can also
   improve performance by reducing the congestion on the ACK channel.
   These mechanisms should be employed in asymmetric networks.

   Several researchers have introduced strategies for coping with
   bandwidth asymmetry.  These mechanisms generally attempt to reduce
   the number of ACKs being transmitted over the low bandwidth channel
   by limiting the ACK frequency or filtering out ACKs at an
   intermediate router [BPK98].  While these solutions mitigate the
   performance problems caused by asymmetric subnetworks they do have
   some cost and therefore, as suggested above, bandwidth asymmetry
   should be minimized whenever possible when designing subnetworks.

Buffering, flow & congestion control

   [atm dropping individual cells in a packet means the entire packet
   must be dropped unless EPD/PPD is used]


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   User data compression is a function that can usually be omitted at
   the subnetwork layer. 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 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 [RFC2393] supports DEFLATE [RFC2394] and LZS
   [RFC2395]).  In any of these cases, compression in the subnetwork is
   of no benefit.

   The subnetwork may also process data that has been encrypted at the
   application protocol layer (OpenPGP [RFC2440] or S/MIME
   [RFCs-2630-2634]), the transport layer (SSL, TLS [RFC2246]), or the
   IP layer (IPSEC ESP [RFC2406]). Ciphers generate random-looking bit
   streams lacking any patterns that can be exploited by a compression

   If a subnetwork decides to implement user data compression, it must
   detect when the data is encrypted or already compressed and transmit
   it without further compression. This is important because most
   compression algorithms increase the size of encrypted data or data
   that has already been compressed.

   In contrast to user data compression, subnetworks that operate at low
   speed or with small packet size limits are encouraged to compress IP
   and transport-level headers (TCP and UDP). 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. This is
   especially beneficial for small, latency-sensitive packets, such as
   in interactive sessions.

   Designers should consider the effect of the subnetwork error rate on
   performance when considering 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 about data corruption and
   subsequent loss of synchronization between compressor and
   decompressor. It relies instead on TCP retransmission to
   resynchronize the decompressor.  After a packet is lost, the
   decompressor must discard every subsequent packet, even if the

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   subnetwork makes no further errors, until the sending TCP retransmits
   to resynchronize 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.

   Alternative 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.

Packet Reordering

   The Internet architecture does not guarantee that packets will arrive
   in the same order in which they were originally transmitted, and
   transport protocols like TCP must take this into account.  However,
   we recommend that subnetworks not gratuitously deliver packets out of
   sequence.  Since 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,APS99] to resend the segment.  In
   addition, the congestion window is reduced by half, effectively
   halving TCP's sending rate.  If a subnetwork badly re-orders segments
   such that three duplicate ACKs are generated the TCP sender
   needlessly reduces the congestion window, and therefore performance.


   [best provided at a higher layer, for performance and flexibility
   reasons, but some subnet mobility can be a convenience as long as
   it's not too inefficient with routing]


   Similar to the case of broadcast and discovery, multicast is more
   efficient on shared links where it is supported natively. Native
   multicast support requires a reasonable number (?? - over 10, under
   1000?) of separate link-layer broadcast addresses. One such address
   SHOULD be reserved for native link broadcast; other addresses SHOULD
   be provided support separate multicast groups (and there SHOULD be at
   least 10?? such addresses).

   The other criteria for native multicast is a link-layer filter, which
   can select individual or sets of broadcast addresses. Such link
   filters avoid having every host parse every multicast message in the

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   driver; a host receives, at the network layer, only those packets
   that pass its configured link filters. A shared link SHOULD support
   multiple, programmable link filters, to support efficient native

   [Multicasting can be simulated over unicast subnets by sending
   multiple copies of packets, but this is wasteful. If the subnet can
   support native multicasting in an efficient way, it should do so]

Broadcasting and Discovery

   Link layers fall into two categories: point-to-point and shared link.
   A point-to-point link has exactly two endpoint components (hosts or
   gateways); a shared link has more than two, either on an inherently
   broadcast media (e.g., Ethernet, radio) or on a switching layer
   hidden from the network layer (switched Ethernet, Myrinet, ATM).

   There are a number of Internet protocols which make use of link layer
   broadcast capabilities. These include link layer address lookup
   (ARP), auto-configuration (RARP, BOOTP, DHCP), and routing (RIP).
   These protocols require broadcast-capable links. Shared links SHOULD
   support native, link layer subnet broadcast.

   The lack of broadcast can impede the performance of these protocols,
   or in some cases render them inoperable. ARP-like link address lookup
   can be provided by a centralized database, rather than owner response
   to broadcast queries. This comes at the expense of potentially higher
   response latency and the need for explicit knowledge of the ARP
   server address (no automatic ARP discovery).

   For other protocols, if a link does not support broadcast, the
   protocol is inoperable. This is the case for DHCP, for example.


   [what is proper division between routing at the Internet layer and
   routing in the subnet? Is it useful or helpful to Internet routing to
   have subnetworks that provide their own internal routing?]


   [Security mechanisms should be placed as close as possible to the
   entities that they protect. E.g., mechanisms that protect host
   computers or users should be implemented at the higher layers and
   operate on an end-to-end basis under control of the users. This makes
   subnet security mechanisms largely redundant unless they are to
   protect the subnet itself, e.g., against unauthorized use.]

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INTERNET DRAFT                                          October 22, 1999


   References of the form RFCnnnn are Internet Request for Comments
   (RFC) documents available online at

   [APS99] Mark Allman, Vern Paxson, W. Richard Stevens.  TCP Congestion
   Control, April 1999.  RFC 2581.

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

   [Jac90] Van Jacobson.  Modified TCP Congestion Avoidance Algorithm.
   Email to the end2end-interest mailing list, April 1990.  URL:

   [SRC81] Jerome H. Saltzer, David P. Reed and David 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. Also scheduled to be reprinted in
   Amit Bhargava, editor. Integrated broadband networks.  Artech House,
   Boston, 1991. ISBN 0-89006-483-0.

   [RFC791] Jon Postel.  "Internet Protocol". September 1981.

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

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

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

   [RFC1577] M. Laubach.  "Classical IP and ARP over ATM". January 1994.

   [RFC1661] W. Simpson. "he Point-to-Point Protocol (PPP)". July 1994.

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

   [RFC2364] G. Gross et al. "PPP Over AAL5". July 1998.

   [RFC2393] A. Shacham et al. "IP Payload Compression Protocol
   (IPComp)". December 1998.

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INTERNET DRAFT                                          October 22, 1999

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

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

   [RFC2440] J. Callas et al. "OpenPGP Message Format". November 1998.

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

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

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

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

   [RFC2406] S. Kent, R. Atkinson. "P Encapsulating Security Payload
   (ESP)". November 1998.

   [RFC2616] R. Fielding et al. "Hypertext Transfer Protocol --
   HTTP/1.1". June 1999.

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

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

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

   [OKM96] T. Ott, J.H.B. Kemperman, M. Mathis, The Stationary Behavior
   of Ideal TCP Congestion Avoidance.

   [RED93] S. Floyd, V. Jacobson, "Random Early Detection gateways for
   Congestion Avoidance", IEEE/ACM Transactions in Networking, V.1 N.4,
   August 1993,

   [Stevens94] R. Stevens, "TCP/IP Illustrated, Volume 1," Addison-
   Wesley, 1994 (section 2.10).

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Security Considerations

   [comment here]

Authors'  Addresses:

   Phil Karn (
   Aaron Falk (
   Joe Touch (
   Marie-Jose Montpetit (
   Jamshid Mahdavi (
   Gabriel Montenegro (

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