





Network Working Group                                  M. Allman, Editor
Request for Comments: 2760   NASA Glenn Research Center/BBN Technologies
Category: Informational                                       S. Dawkins
                                                                  Nortel
                                                               D. Glover
                                                               J. Griner
                                                                 D. Tran
                                              NASA Glenn Research Center
                                                            T. Henderson
                                    University of California at Berkeley
                                                            J. Heidemann
                                                                J. Touch
                                   University of Southern California/ISI
                                                                H. Kruse
                                                            S. Ostermann
                                                         Ohio University
                                                                K. Scott
                                                   The MITRE Corporation
                                                                J. Semke
                                        Pittsburgh Supercomputing Center
                                                           February 2000


               Ongoing TCP Research Related to Satellites


Status of this Memo

   This memo provides information for the Internet community.  It does
   not specify an Internet standard of any kind.  Distribution of this
   memo is unlimited.

Copyright Notice

   Copyright (C) The Internet Society (2000).  All Rights Reserved.

Abstract

   This document outlines possible TCP enhancements that may allow TCP
   to better utilize the available bandwidth provided by networks
   containing satellite links.  The algorithms and mechanisms outlined
   have not been judged to be mature enough to be recommended by the
   IETF.  The goal of this document is to educate researchers as to the
   current work and progress being done in TCP research related to
   satellite networks.






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Table of Contents

   1         Introduction. . . . . . . . . . . . . . . . . . . .  2
   2         Satellite Architectures . . . . . . . . . . . . . .  3
   2.1       Asymmetric Satellite Networks . . . . . . . . . . .  3
   2.2       Satellite Link as Last Hop. . . . . . . . . . . . .  3
   2.3       Hybrid Satellite Networks     . . . . . . . . . . .  4
   2.4       Point-to-Point Satellite Networks . . . . . . . . .  4
   2.5       Multiple Satellite Hops . . . . . . . . . . . . . .  4
   3         Mitigations . . . . . . . . . . . . . . . . . . . .  4
   3.1       TCP For Transactions. . . . . . . . . . . . . . . .  4
   3.2       Slow Start. . . . . . . . . . . . . . . . . . . . .  5
   3.2.1     Larger Initial Window . . . . . . . . . . . . . . .  6
   3.2.2     Byte Counting . . . . . . . . . . . . . . . . . . .  7
   3.2.3     Delayed ACKs After Slow Start . . . . . . . . . . .  9
   3.2.4     Terminating Slow Start. . . . . . . . . . . . . . . 11
   3.3       Loss Recovery . . . . . . . . . . . . . . . . . . . 12
   3.3.1     Non-SACK Based Mechanisms . . . . . . . . . . . . . 12
   3.3.2     SACK Based Mechanisms . . . . . . . . . . . . . . . 13
   3.3.3     Explicit Congestion Notification. . . . . . . . . . 16
   3.3.4     Detecting Corruption Loss . . . . . . . . . . . . . 18
   3.4       Congestion Avoidance. . . . . . . . . . . . . . . . 21
   3.5       Multiple Data Connections . . . . . . . . . . . . . 22
   3.6       Pacing TCP Segments . . . . . . . . . . . . . . . . 24
   3.7       TCP Header Compression. . . . . . . . . . . . . . . 26
   3.8       Sharing TCP State Among Similar Connections . . . . 29
   3.9       ACK Congestion Control. . . . . . . . . . . . . . . 32
   3.10      ACK Filtering . . . . . . . . . . . . . . . . . . . 34
   4         Conclusions . . . . . . . . . . . . . . . . . . . . 36
   5         Security Considerations . . . . . . . . . . . . . . 36
   6         Acknowledgments . . . . . . . . . . . . . . . . . . 37
   7         References. . . . . . . . . . . . . . . . . . . . . 37
   8         Authors' Addresses. . . . . . . . . . . . . . . . . 43
   9         Full Copyright Statement. . . . . . . . . . . . . . 46

1   Introduction

   This document outlines mechanisms that may help the Transmission
   Control Protocol (TCP) [Pos81] better utilize the bandwidth provided
   by long-delay satellite environments.  These mechanisms may also help
   in other environments or for other protocols.  The proposals outlined
   in this document are currently being studied throughout the research
   community.  Therefore, these mechanisms are not mature enough to be
   recommended for wide-spread use by the IETF.  However, some of these
   mechanisms may be safely used today.  It is hoped that this document
   will stimulate further study into the described mechanisms.  If, at





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   some point, the mechanisms discussed in this memo prove to be safe
   and appropriate to be recommended for general use, the appropriate
   IETF documents will be written.

   It should be noted that non-TCP mechanisms that help performance over
   satellite links do exist (e.g., application-level changes, queueing
   disciplines, etc.).  However, outlining these non-TCP mitigations is
   beyond the scope of this document and therefore is left as future
   work.  Additionally, there are a number of mitigations to TCP's
   performance problems that involve very active intervention by
   gateways along the end-to-end path from the sender to the receiver.
   Documenting the pros and cons of such solutions is also left as
   future work.

2   Satellite Architectures

   Specific characteristics of satellite links and the impact these
   characteristics have on TCP are presented in RFC 2488 [AGS99].  This
   section discusses several possible topologies where satellite links
   may be integrated into the global Internet.  The mitigation outlined
   in section 3 will include a discussion of which environment the
   mechanism is expected to benefit.

2.1 Asymmetric Satellite Networks

   Some satellite networks exhibit a bandwidth asymmetry, a larger data
   rate in one direction than the reverse direction, because of limits
   on the transmission power and the antenna size at one end of the
   link.  Meanwhile, some other satellite systems are unidirectional and
   use a non-satellite return path (such as a dialup modem link).  The
   nature of most TCP traffic is asymmetric with data flowing in one
   direction and acknowledgments in opposite direction.  However, the
   term asymmetric in this document refers to different physical
   capacities in the forward and return links.  Asymmetry has been shown
   to be a problem for TCP [BPK97,BPK98].

2.2 Satellite Link as Last Hop

   Satellite links that provide service directly to end users, as
   opposed to satellite links located in the middle of a network, may
   allow for specialized design of protocols used over the last hop.
   Some satellite providers use the satellite link as a shared high
   speed downlink to users with a lower speed, non-shared terrestrial
   link that is used as a return link for requests and acknowledgments.
   Many times this creates an asymmetric network, as discussed above.






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2.3 Hybrid Satellite Networks

   In the more general case, satellite links may be located at any point
   in the network topology.  In this case, the satellite link acts as
   just another link between two gateways.  In this environment, a given
   connection may be sent over terrestrial links (including terrestrial
   wireless), as well as satellite links.  On the other hand, a
   connection could also travel over only the terrestrial network or
   only over the satellite portion of the network.

2.4 Point-to-Point Satellite Networks

   In point-to-point satellite networks, the only hop in the network is
   over the satellite link.  This pure satellite environment exhibits
   only the problems associated with the satellite links, as outlined in
   [AGS99].  Since this is a private network, some mitigations that are
   not appropriate for shared networks can be considered.

2.5 Multiple Satellite Hops

   In some situations, network traffic may traverse multiple satellite
   hops between the source and the destination.  Such an environment
   aggravates the satellite characteristics described in [AGS99].

3   Mitigations

   The following sections will discuss various techniques for mitigating
   the problems TCP faces in the satellite environment.  Each of the
   following sections will be organized as follows: First, each
   mitigation will be briefly outlined.  Next, research work involving
   the mechanism in question will be briefly discussed.  Next the
   implementation issues of the mechanism will be presented (including
   whether or not the particular mechanism presents any dangers to
   shared networks).  Then a discussion of the mechanism's potential
   with regard to the topologies outlined above is given.  Finally, the
   relationships and possible interactions with other TCP mechanisms are
   outlined.  The reader is expected to be familiar with the TCP
   terminology used in [AGS99].

3.1 TCP For Transactions

3.1.1 Mitigation Description

   TCP uses a three-way handshake to setup a connection between two
   hosts [Pos81].  This connection setup requires 1-1.5 round-trip times
   (RTTs), depending upon whether the data sender started the connection
   actively or passively.  This startup time can be eliminated by using
   TCP extensions for transactions (T/TCP) [Bra94].  After the first



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   connection between a pair of hosts is established, T/TCP is able to
   bypass the three-way handshake, allowing the data sender to begin
   transmitting data in the first segment sent (along with the SYN).
   This is especially helpful for short request/response traffic, as it
   saves a potentially long setup phase when no useful data is being
   transmitted.

3.1.2 Research

   T/TCP is outlined and analyzed in [Bra92,Bra94].

3.1.3 Implementation Issues

   T/TCP requires changes in the TCP stacks of both the data sender and
   the data receiver.  While T/TCP is safe to implement in shared
   networks from a congestion control perspective, several security
   implications of sending data in the first data segment have been
   identified [ddKI99].

3.1.4 Topology Considerations

   It is expected that T/TCP will be equally beneficial in all
   environments outlined in section 2.

3.1.5 Possible Interaction and Relationships with Other Research

   T/TCP allows data transfer to start more rapidly, much like using a
   larger initial congestion window (see section 3.2.1), delayed ACKs
   after slow start (section 3.2.3) or byte counting (section 3.2.2).

3.2 Slow Start

   The slow start algorithm is used to gradually increase the size of
   TCP's congestion window (cwnd) [Jac88,Ste97,APS99].  The algorithm is
   an important safe-guard against transmitting an inappropriate amount
   of data into the network when the connection starts up.  However,
   slow start can also waste available network capacity, especially in
   long-delay networks [All97a,Hay97].  Slow start is particularly
   inefficient for transfers that are short compared to the
   delay*bandwidth product of the network (e.g., WWW transfers).

   Delayed ACKs are another source of wasted capacity during the slow
   start phase.  RFC 1122 [Bra89] suggests data receivers refrain from
   ACKing every incoming data segment.  However, every second full-sized
   segment should be ACKed.  If a second full-sized segment does not
   arrive within a given timeout, an ACK must be generated (this timeout
   cannot exceed 500 ms).  Since the data sender increases the size of
   cwnd based on the number of arriving ACKs, reducing the number of



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   ACKs slows the cwnd growth rate.  In addition, when TCP starts
   sending, it sends 1 segment.  When using delayed ACKs a second
   segment must arrive before an ACK is sent.  Therefore, the receiver
   is always forced to wait for the delayed ACK timer to expire before
   ACKing the first segment, which also increases the transfer time.

   Several proposals have suggested ways to make slow start less time
   consuming.  These proposals are briefly outlined below and references
   to the research work given.

3.2.1 Larger Initial Window

3.2.1.1 Mitigation Description

   One method that will reduce the amount of time required by slow start
   (and therefore, the amount of wasted capacity) is to increase the
   initial value of cwnd.  An experimental TCP extension outlined in
   [AFP98] allows the initial size of cwnd to be increased from 1
   segment to that given in equation (1).

               min (4*MSS, max (2*MSS, 4380 bytes))               (1)

   By increasing the initial value of cwnd, more packets are sent during
   the first RTT of data transmission, which will trigger more ACKs,
   allowing the congestion window to open more rapidly.  In addition, by
   sending at least 2 segments initially, the first segment does not
   need to wait for the delayed ACK timer to expire as is the case when
   the initial size of cwnd is 1 segment (as discussed above).
   Therefore, the value of cwnd given in equation 1 saves up to 3 RTTs
   and a delayed ACK timeout when compared to an initial cwnd of 1
   segment.

   Also, we note that RFC 2581 [APS99], a standards-track document,
   allows a TCP to use an initial cwnd of up to 2 segments.  This change
   is highly recommended for satellite networks.

3.2.1.2 Research

   Several researchers have studied the use of a larger initial window
   in various environments.  [Nic97] and [KAGT98] show a reduction in
   WWW page transfer time over hybrid fiber coax (HFC) and satellite
   links respectively.  Furthermore, it has been shown that using an
   initial cwnd of 4 segments does not negatively impact overall
   performance over dialup modem links with a small number of buffers
   [SP98].  [AHO98] shows an improvement in transfer time for 16 KB
   files across the Internet and dialup modem links when using a larger
   initial value for cwnd.  However, a slight increase in dropped




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   segments was also shown.  Finally, [PN98] shows improved transfer
   time for WWW traffic in simulations with competing traffic, in
   addition to a small increase in the drop rate.

3.2.1.3 Implementation Issues

   The use of a larger initial cwnd value requires changes to the
   sender's TCP stack.  Using an initial congestion window of 2 segments
   is allowed by RFC 2581 [APS99].  Using an initial congestion window
   of 3 or 4 segments is not expected to present any danger of
   congestion collapse [AFP98], however may degrade performance in some
   networks.

3.2.1.4 Topology Considerations

   It is expected that the use of a large initial window would be
   equally beneficial to all network architectures outlined in section
   2.

3.2.1.5 Possible Interaction and Relationships with Other Research

   Using a fixed larger initial congestion window decreases the impact
   of a long RTT on transfer time (especially for short transfers) at
   the cost of bursting data into a network with unknown conditions.  A
   mechanism that mitigates bursts may make the use of a larger initial
   congestion window more appropriate (e.g., limiting the size of line-
   rate bursts [FF96] or pacing the segments in a burst [VH97a]).

   Also, using delayed ACKs only after slow start (as outlined in
   section 3.2.3) offers an alternative way to immediately ACK the first
   segment of a transfer and open the congestion window more rapidly.
   Finally, using some form of TCP state sharing among a number of
   connections (as discussed in 3.8) may provide an alternative to using
   a fixed larger initial window.

3.2.2 Byte Counting

3.2.2.1 Mitigation Description

   As discussed above, the wide-spread use of delayed ACKs increases the
   time needed by a TCP sender to increase the size of the congestion
   window during slow start.  This is especially harmful to flows
   traversing long-delay GEO satellite links.  One mechanism that has
   been suggested to mitigate the problems caused by delayed ACKs is the
   use of "byte counting", rather than standard ACK counting
   [All97a,All98].  Using standard ACK counting, the congestion window
   is increased by 1 segment for each ACK received during slow start.
   However, using byte counting the congestion window increase is based



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   on the number of previously unacknowledged bytes covered by each
   incoming ACK, rather than on the number of ACKs received.  This makes
   the increase relative to the amount of data transmitted, rather than
   being dependent on the ACK interval used by the receiver.

   Two forms of byte counting are studied in [All98].  The first is
   unlimited byte counting (UBC).  This mechanism simply uses the number
   of previously unacknowledged bytes to increase the congestion window
   each time an ACK arrives.  The second form is limited byte counting
   (LBC).  LBC limits the amount of cwnd increase to 2 segments.  This
   limit throttles the size of the burst of data sent in response to a
   "stretch ACK" [Pax97].  Stretch ACKs are acknowledgments that cover
   more than 2 segments of previously unacknowledged data.  Stretch ACKs
   can occur by design [Joh95] (although this is not standard), due to
   implementation bugs [All97b,PADHV99] or due to ACK loss.  [All98]
   shows that LBC prevents large line-rate bursts when compared to UBC,
   and therefore offers fewer dropped segments and better performance.
   In addition, UBC causes large bursts during slow start based loss
   recovery due to the large cumulative ACKs that can arrive during loss
   recovery.  The behavior of UBC during loss recovery can cause large
   decreases in performance and [All98] strongly recommends UBC not be
   deployed without further study into mitigating the large bursts.

   Note: The standards track RFC 2581 [APS99] allows a TCP to use byte
   counting to increase cwnd during congestion avoidance, however not
   during slow start.

3.2.2.2 Research

   Using byte counting, as opposed to standard ACK counting, has been
   shown to reduce the amount of time needed to increase the value of
   cwnd to an appropriate size in satellite networks [All97a].  In
   addition, [All98] presents a simulation comparison of byte counting
   and the standard cwnd increase algorithm in uncongested networks and
   networks with competing traffic.  This study found that the limited
   form of byte counting outlined above can improve performance, while
   also increasing the drop rate slightly.

   [BPK97,BPK98] also investigated unlimited byte counting in
   conjunction with various ACK filtering algorithms (discussed in
   section 3.10) in asymmetric networks.










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3.2.2.3 Implementation Issues

   Changing from ACK counting to byte counting requires changes to the
   data sender's TCP stack.  Byte counting violates the algorithm for
   increasing the congestion window outlined in RFC 2581 [APS99] (by
   making congestion window growth more aggressive during slow start)
   and therefore should not be used in shared networks.

3.2.2.4 Topology Considerations

   It has been suggested by some (and roundly criticized by others) that
   byte counting will allow TCP to provide uniform cwnd increase,
   regardless of the ACKing behavior of the receiver.  In addition, byte
   counting also mitigates the retarded window growth provided by
   receivers that generate stretch ACKs because of the capacity of the
   return link, as discussed in [BPK97,BPK98].  Therefore, this change
   is expected to be especially beneficial to asymmetric networks.

3.2.2.5 Possible Interaction and Relationships with Other Research

   Unlimited byte counting should not be used without a method to
   mitigate the potentially large line-rate bursts the algorithm can
   cause.  Also, LBC may send bursts that are too large for the given
   network conditions.  In this case, LBC may also benefit from some
   algorithm that would lessen the impact of line-rate bursts of
   segments.  Also note that using delayed ACKs only after slow start
   (as outlined in section 3.2.3) negates the limited byte counting
   algorithm because each ACK covers only one segment during slow start.
   Therefore, both ACK counting and byte counting yield the same
   increase in the congestion window at this point (in the first RTT).

3.2.3 Delayed ACKs After Slow Start

3.2.3.1 Mitigation Description

   As discussed above, TCP senders use the number of incoming ACKs to
   increase the congestion window during slow start.  And, since delayed
   ACKs reduce the number of ACKs returned by the receiver by roughly
   half, the rate of growth of the congestion window is reduced.  One
   proposed solution to this problem is to use delayed ACKs only after
   the slow start (DAASS) phase.  This provides more ACKs while TCP is
   aggressively increasing the congestion window and less ACKs while TCP
   is in steady state, which conserves network resources.








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3.2.3.2 Research

   [All98] shows that in simulation, using delayed ACKs after slow start
   (DAASS) improves transfer time when compared to a receiver that
   always generates delayed ACKs.  However, DAASS also slightly
   increases the loss rate due to the increased rate of cwnd growth.

3.2.3.3 Implementation Issues

   The major problem with DAASS is in the implementation.  The receiver
   has to somehow know when the sender is using the slow start
   algorithm.  The receiver could implement a heuristic that attempts to
   watch the change in the amount of data being received and change the
   ACKing behavior accordingly.  Or, the sender could send a message (a
   flipped bit in the TCP header, perhaps) indicating that it was using
   slow start.  The implementation of DAASS is, therefore, an open
   issue.

   Using DAASS does not violate the TCP congestion control specification
   [APS99].  However, the standards (RFC 2581 [APS99]) currently
   recommend using delayed acknowledgments and DAASS goes (partially)
   against this recommendation.

3.2.3.4 Topology Considerations

   DAASS should work equally well in all scenarios presented in section
   2.  However, in asymmetric networks it may aggravate ACK congestion
   in the return link, due to the increased number of ACKs (see sections
   3.9 and 3.10 for a more detailed discussion of ACK congestion).

3.2.3.5 Possible Interaction and Relationships with Other Research

   DAASS has several possible interactions with other proposals made in
   the research community.  DAASS can aggravate congestion on the path
   between the data receiver and the data sender due to the increased
   number of returning acknowledgments.  This can have an especially
   adverse effect on asymmetric networks that are prone to experiencing
   ACK congestion.  As outlined in sections 3.9 and 3.10, several
   mitigations have been proposed to reduce the number of ACKs that are
   passed over a low-bandwidth return link.  Using DAASS will increase
   the number of ACKs sent by the receiver.  The interaction between
   DAASS and the methods for reducing the number of ACKs is an open
   research question.  Also, as noted in section 3.2.1.5 above, DAASS
   provides some of the same benefits as using a larger initial
   congestion window and therefore it may not be desirable to use both
   mechanisms together.  However, this remains an open question.
   Finally, DAASS and limited byte counting are both used to increase




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   the rate at which the congestion window is opened.  The DAASS
   algorithm substantially reduces the impact limited byte counting has
   on the rate of congestion window increase.

3.2.4 Terminating Slow Start

3.2.4.1 Mitigation Description

   The initial slow start phase is used by TCP to determine an
   appropriate congestion window size for the given network conditions
   [Jac88].  Slow start is terminated when TCP detects congestion, or
   when the size of cwnd reaches the size of the receiver's advertised
   window.  Slow start is also terminated if cwnd grows beyond a certain
   size.  The threshold at which TCP ends slow start and begins using
   the congestion avoidance algorithm is called "ssthresh" [Jac88].  In
   most implementations, the initial value for ssthresh is the
   receiver's advertised window.  During slow start, TCP roughly doubles
   the size of cwnd every RTT and therefore can overwhelm the network
   with at most twice as many segments as the network can handle.  By
   setting ssthresh to a value less than the receiver's advertised
   window initially, the sender may avoid overwhelming the network with
   twice the appropriate number of segments.  Hoe [Hoe96] proposes using
   the packet-pair algorithm [Kes91] and the measured RTT to determine a
   more appropriate value for ssthresh.  The algorithm observes the
   spacing between the first few returning ACKs to determine the
   bandwidth of the bottleneck link.  Together with the measured RTT,
   the delay*bandwidth product is determined and ssthresh is set to this
   value.  When TCP's cwnd reaches this reduced ssthresh, slow start is
   terminated and transmission continues using congestion avoidance,
   which is a more conservative algorithm for increasing the size of the
   congestion window.

3.2.4.2 Research

   It has been shown that estimating ssthresh can improve performance
   and decrease packet loss in simulations [Hoe96].  However, obtaining
   an accurate estimate of the available bandwidth in a dynamic network
   is very challenging, especially attempting to do so on the sending
   side of the TCP connection [AP99].  Therefore, before this mechanism
   is widely deployed, bandwidth estimation must be studied in a more
   detail.

3.2.4.3 Implementation Issues

   As outlined in [Hoe96], estimating ssthresh requires changes to the
   data sender's TCP stack.  As suggested in [AP99], bandwidth estimates
   may be more accurate when taken by the TCP receiver, and therefore
   both sender and receiver changes would be required.  Estimating



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   ssthresh is safe to implement in production networks from a
   congestion control perspective, as it can only make TCP more
   conservative than outlined in RFC 2581 [APS99] (assuming the TCP
   implementation is using an initial ssthresh of infinity as allowed by
   [APS99]).

3.2.4.4 Topology Considerations

   It is expected that this mechanism will work equally well in all
   symmetric topologies outlined in section 2.  However, asymmetric
   links pose a special problem, as the rate of the returning ACKs may
   not be the bottleneck bandwidth in the forward direction.  This can
   lead to the sender setting ssthresh too low.  Premature termination
   of slow start can hurt performance, as congestion avoidance opens
   cwnd more conservatively.  Receiver-based bandwidth estimators do not
   suffer from this problem.

3.2.4.5 Possible Interaction and Relationships with Other Research

   Terminating slow start at the right time is useful to avoid multiple
   dropped segments.  However, using a selective acknowledgment-based
   loss recovery scheme (as outlined in section 3.3.2) can drastically
   improve TCP's ability to quickly recover from multiple lost segments
   Therefore, it may not be as important to terminate slow start before
   a large loss event occurs.  [AP99] shows that using delayed
   acknowledgments [Bra89] reduces the effectiveness of sender-side
   bandwidth estimation.  Therefore, using delayed ACKs only during slow
   start (as outlined in section 3.2.3) may make bandwidth estimation
   more feasible.

3.3 Loss Recovery

3.3.1 Non-SACK Based Mechanisms

3.3.1.1 Mitigation Description

   Several similar algorithms have been developed and studied that
   improve TCP's ability to recover from multiple lost segments in a
   window of data without relying on the (often long) retransmission
   timeout.  These sender-side algorithms, known as NewReno TCP, do not
   depend on the availability of selective acknowledgments (SACKs)
   [MMFR96].

   These algorithms generally work by updating the fast recovery
   algorithm to use information provided by "partial ACKs" to trigger
   retransmissions.  A partial ACK covers some new data, but not all
   data outstanding when a particular loss event starts.  For instance,
   consider the case when segment N is retransmitted using the fast



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   retransmit algorithm and segment M is the last segment sent when
   segment N is resent.  If segment N is the only segment lost, the ACK
   elicited by the retransmission of segment N would be for segment M.
   If, however, segment N+1 was also lost, the ACK elicited by the
   retransmission of segment N will be N+1.  This can be taken as an
   indication that segment N+1 was lost and used to trigger a
   retransmission.

3.3.1.2 Research

   Hoe [Hoe95,Hoe96] introduced the idea of using partial ACKs to
   trigger retransmissions and showed that doing so could improve
   performance.  [FF96] shows that in some cases using partial ACKs to
   trigger retransmissions reduces the time required to recover from
   multiple lost segments.  However, [FF96] also shows that in some
   cases (many lost segments) relying on the RTO timer can improve
   performance over simply using partial ACKs to trigger all
   retransmissions.  [HK99] shows that using partial ACKs to trigger
   retransmissions, in conjunction with SACK, improves performance when
   compared to TCP using fast retransmit/fast recovery in a satellite
   environment.  Finally, [FH99] describes several slightly different
   variants of NewReno.

3.3.1.3 Implementation Issues

   Implementing these fast recovery enhancements requires changes to the
   sender-side TCP stack.  These changes can safely be implemented in
   production networks and are allowed by RFC 2581 [APS99].

3.3.1.4 Topology Considerations

   It is expected that these changes will work well in all environments
   outlined in section 2.

3.3.1.5 Possible Interaction and Relationships with Other Research

   See section 3.3.2.2.5.

3.3.2 SACK Based Mechanisms

3.3.2.1 Fast Recovery with SACK

3.3.2.1.1 Mitigation Description

   Fall and Floyd [FF96] describe a conservative extension to the fast
   recovery algorithm that takes into account information provided by
   selective acknowledgments (SACKs) [MMFR96] sent by the receiver.  The
   algorithm starts after fast retransmit triggers the resending of a



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   segment.  As with fast retransmit, the algorithm cuts cwnd in half
   when a loss is detected.  The algorithm keeps a variable called
   "pipe", which is an estimate of the number of outstanding segments in
   the network.  The pipe variable is decremented by 1 segment for each
   duplicate ACK that arrives with new SACK information.  The pipe
   variable is incremented by 1 for each new or retransmitted segment
   sent.  A segment may be sent when the value of pipe is less than cwnd
   (this segment is either a retransmission per the SACK information or
   a new segment if the SACK information indicates that no more
   retransmits are needed).

   This algorithm generally allows TCP to recover from multiple segment
   losses in a window of data within one RTT of loss detection.  Like
   the forward acknowledgment (FACK) algorithm described below, the SACK
   information allows the pipe algorithm to decouple the choice of when
   to send a segment from the choice of what segment to send.

   [APS99] allows the use of this algorithm, as it is consistent with
   the spirit of the fast recovery algorithm.

3.3.2.1.2 Research

   [FF96] shows that the above described SACK algorithm performs better
   than several non-SACK based recovery algorithms when 1--4 segments
   are lost from a window of data.  [AHKO97] shows that the algorithm
   improves performance over satellite links.  Hayes [Hay97] shows the
   in certain circumstances, the SACK algorithm can hurt performance by
   generating a large line-rate burst of data at the end of loss
   recovery, which causes further loss.

3.3.2.1.3 Implementation Issues

   This algorithm is implemented in the sender's TCP stack.  However, it
   relies on SACK information generated by the receiver.  This algorithm
   is safe for shared networks and is allowed by RFC 2581 [APS99].

3.3.2.1.4 Topology Considerations

   It is expected that the pipe algorithm will work equally well in all
   scenarios presented in section 2.

3.3.2.1.5 Possible Interaction and Relationships with Other Research

   See section 3.3.2.2.5.







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3.3.2.2 Forward Acknowledgments

3.3.2.2.1 Mitigation Description

   The Forward Acknowledgment (FACK) algorithm [MM96a,MM96b] was
   developed to improve TCP congestion control during loss recovery.
   FACK uses TCP SACK options to glean additional information about the
   congestion state, adding more precise control to the injection of
   data into the network during recovery.  FACK decouples the congestion
   control algorithms from the data recovery algorithms to provide a
   simple and direct way to use SACK to improve congestion control.  Due
   to the separation of these two algorithms, new data may be sent
   during recovery to sustain TCP's self-clock when there is no further
   data to retransmit.

   The most recent version of FACK is Rate-Halving [MM96b], in which one
   packet is sent for every two ACKs received during recovery.
   Transmitting a segment for every-other ACK has the result of reducing
   the congestion window in one round trip to half of the number of
   packets that were successfully handled by the network (so when cwnd
   is too large by more than a factor of two it still gets reduced to
   half of what the network can sustain).  Another important aspect of
   FACK with Rate-Halving is that it sustains the ACK self-clock during
   recovery because transmitting a packet for every-other ACK does not
   require half a cwnd of data to drain from the network before
   transmitting, as required by the fast recovery algorithm
   [Ste97,APS99].

   In addition, the FACK with Rate-Halving implementation provides
   Thresholded Retransmission to each lost segment.  "Tcprexmtthresh" is
   the number of duplicate ACKs required by TCP to trigger a fast
   retransmit and enter recovery.  FACK applies thresholded
   retransmission to all segments by waiting until tcprexmtthresh SACK
   blocks indicate that a given segment is missing before resending the
   segment.  This allows reasonable behavior on links that reorder
   segments.  As described above, FACK sends a segment for every second
   ACK received during recovery.  New segments are transmitted except
   when tcprexmtthresh SACK blocks have been observed for a dropped
   segment, at which point the dropped segment is retransmitted.

   [APS99] allows the use of this algorithm, as it is consistent with
   the spirit of the fast recovery algorithm.

3.3.2.2.2 Research

   The original FACK algorithm is outlined in [MM96a].  The algorithm
   was later enhanced to include Rate-Halving [MM96b].  The real-world
   performance of FACK with Rate-Halving was shown to be much closer to



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   the theoretical maximum for TCP than either TCP Reno or the SACK-
   based extensions to fast recovery outlined in section 3.3.2.1
   [MSMO97].

3.3.2.2.3 Implementation Issues

   In order to use FACK, the sender's TCP stack must be modified.  In
   addition, the receiver must be able to generate SACK options to
   obtain the full benefit of using FACK.  The FACK algorithm is safe
   for shared networks and is allowed by RFC 2581 [APS99].

3.3.2.2.4 Topology Considerations

   FACK is expected to improve performance in all environments outlined
   in section 2.  Since it is better able to sustain its self-clock than
   TCP Reno, it may be considerably more attractive over long delay
   paths.

3.3.2.2.5 Possible Interaction and Relationships with Other Research

   Both SACK based loss recovery algorithms described above (the fast
   recovery enhancement and the FACK algorithm) are similar in that they
   attempt to effectively repair multiple lost segments from a window of
   data.  Which of the SACK-based loss recovery algorithms to use is
   still an open research question.  In addition, these algorithms are
   similar to the non-SACK NewReno algorithm described in section 3.3.1,
   in that they attempt to recover from multiple lost segments without
   reverting to using the retransmission timer.  As has been shown, the
   above SACK based algorithms are more robust than the NewReno
   algorithm.  However, the SACK algorithm requires a cooperating TCP
   receiver, which the NewReno algorithm does not.  A reasonable TCP
   implementation might include both a SACK-based and a NewReno-based
   loss recovery algorithm such that the sender can use the most
   appropriate loss recovery algorithm based on whether or not the
   receiver supports SACKs.  Finally, both SACK-based and non-SACK-based
   versions of fast recovery have been shown to transmit a large burst
   of data upon leaving loss recovery, in some cases [Hay97].
   Therefore, the algorithms may benefit from some burst suppression
   algorithm.

3.3.3 Explicit Congestion Notification

3.3.3.1 Mitigation Description

   Explicit congestion notification (ECN) allows routers to inform TCP
   senders about imminent congestion without dropping segments.  Two
   major forms of ECN have been studied.  A router employing backward
   ECN (BECN), transmits messages directly to the data originator



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   informing it of congestion.  IP routers can accomplish this with an
   ICMP Source Quench message.  The arrival of a BECN signal may or may
   not mean that a TCP data segment has been dropped, but it is a clear
   indication that the TCP sender should reduce its sending rate (i.e.,
   the value of cwnd).  The second major form of congestion notification
   is forward ECN (FECN).  FECN routers mark data segments with a
   special tag when congestion is imminent, but forward the data
   segment.  The data receiver then echos the congestion information
   back to the sender in the ACK packet.  A description of a FECN
   mechanism for TCP/IP is given in [RF99].

   As described in [RF99], senders transmit segments with an "ECN-
   Capable Transport" bit set in the IP header of each packet.  If a
   router employing an active queueing strategy, such as Random Early
   Detection (RED) [FJ93,BCC+98], would otherwise drop this segment, an
   "Congestion Experienced" bit in the IP header is set instead.  Upon
   reception, the information is echoed back to TCP senders using a bit
   in the TCP header.  The TCP sender adjusts the congestion window just
   as it would if a segment was dropped.

   The implementation of ECN as specified in [RF99] requires the
   deployment of active queue management mechanisms in the affected
   routers.  This allows the routers to signal congestion by sending TCP
   a small number of "congestion signals" (segment drops or ECN
   messages), rather than discarding a large number of segments, as can
   happen when TCP overwhelms a drop-tail router queue.

   Since satellite networks generally have higher bit-error rates than
   terrestrial networks, determining whether a segment was lost due to
   congestion or corruption may allow TCP to achieve better performance
   in high BER environments than currently possible (due to TCP's
   assumption that all loss is due to congestion).  While not a solution
   to this problem, adding an ECN mechanism to TCP may be a part of a
   mechanism that will help achieve this goal.  See section 3.3.4 for a
   more detailed discussion of differentiating between corruption and
   congestion based losses.

3.3.3.2 Research

   [Flo94] shows that ECN is effective in reducing the segment loss rate
   which yields better performance especially for short and interactive
   TCP connections.  Furthermore, [Flo94] also shows that ECN avoids
   some unnecessary, and costly TCP retransmission timeouts.  Finally,
   [Flo94] also considers some of the advantages and disadvantages of
   various forms of explicit congestion notification.






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3.3.3.3 Implementation Issues

   Deployment of ECN requires changes to the TCP implementation on both
   sender and receiver.  Additionally, deployment of ECN requires
   deployment of some active queue management infrastructure in routers.
   RED is assumed in most ECN discussions, because RED is already
   identifying segments to drop, even before its buffer space is
   exhausted.  ECN simply allows the delivery of "marked" segments while
   still notifying the end nodes that congestion is occurring along the
   path.  ECN is safe (from a congestion control perspective) for shared
   networks, as it maintains the same TCP congestion control principles
   as are used when congestion is detected via segment drops.

3.3.3.4 Topology Considerations

   It is expected that none of the environments outlined in section 2
   will present a bias towards or against ECN traffic.

3.3.3.5 Possible Interaction and Relationships with Other Research

   Note that some form of active queueing is necessary to use ECN (e.g.,
   RED queueing).

3.3.4 Detecting Corruption Loss

   Differentiating between congestion (loss of segments due to router
   buffer overflow or imminent buffer overflow) and corruption (loss of
   segments due to damaged bits) is a difficult problem for TCP.  This
   differentiation is particularly important because the action that TCP
   should take in the two cases is entirely different.  In the case of
   corruption, TCP should merely retransmit the damaged segment as soon
   as its loss is detected; there is no need for TCP to adjust its
   congestion window.  On the other hand, as has been widely discussed
   above, when the TCP sender detects congestion, it should immediately
   reduce its congestion window to avoid making the congestion worse.

   TCP's defined behavior, as motivated by [Jac88,Jac90] and defined in
   [Bra89,Ste97,APS99], is to assume that all loss is due to congestion
   and to trigger the congestion control algorithms, as defined in
   [Ste97,APS99].  The loss may be detected using the fast retransmit
   algorithm, or in the worst case is detected by the expiration of
   TCP's retransmission timer.

   TCP's assumption that loss is due to congestion rather than
   corruption is a conservative mechanism that prevents congestion
   collapse [Jac88,FF98].  Over satellite networks, however, as in many
   wireless environments, loss due to corruption is more common than on
   terrestrial networks.  One common partial solution to this problem is



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   to add Forward Error Correction (FEC) to the data that's sent over
   the satellite/wireless link.  A more complete discussion of the
   benefits of FEC can be found in [AGS99].  However, given that FEC
   does not always work or cannot be universally applied, other
   mechanisms have been studied to attempt to make TCP able to
   differentiate between congestion-based and corruption-based loss.

   TCP segments that have been corrupted are most often dropped by
   intervening routers when link-level checksum mechanisms detect that
   an incoming frame has errors.  Occasionally, a TCP segment containing
   an error may survive without detection until it arrives at the TCP
   receiving host, at which point it will almost always either fail the
   IP header checksum or the TCP checksum and be discarded as in the
   link-level error case.  Unfortunately, in either of these cases, it's
   not generally safe for the node detecting the corruption to return
   information about the corrupt packet to the TCP sender because the
   sending address itself might have been corrupted.

3.3.4.1 Mitigation Description

   Because the probability of link errors on a satellite link is
   relatively greater than on a hardwired link, it is particularly
   important that the TCP sender retransmit these lost segments without
   reducing its congestion window.  Because corrupt segments do not
   indicate congestion, there is no need for the TCP sender to enter a
   congestion avoidance phase, which may waste available bandwidth.
   Simulations performed in [SF98] show a performance improvement when
   TCP can properly differentiate between between corruption and
   congestion of wireless links.

   Perhaps the greatest research challenge in detecting corruption is
   getting TCP (a transport-layer protocol) to receive appropriate
   information from either the network layer (IP) or the link layer.
   Much of the work done to date has involved link-layer mechanisms that
   retransmit damaged segments.  The challenge seems to be to get these
   mechanisms to make repairs in such a way that TCP understands what
   happened and can respond appropriately.

3.3.4.2 Research

   Research into corruption detection to date has focused primarily on
   making the link level detect errors and then perform link-level
   retransmissions.  This work is summarized in [BKVP97,BPSK96].  One of
   the problems with this promising technique is that it causes an
   effective reordering of the segments from the TCP receiver's point of
   view.  As a simple example, if segments A B C D are sent across a
   noisy link and segment B is corrupted, segments C and D may have
   already crossed the link before B can be retransmitted at the link



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   level, causing them to arrive at the TCP receiver in the order A C D
   B.  This segment reordering would cause the TCP receiver to generate
   duplicate ACKs upon the arrival of segments C and D.  If the
   reordering was bad enough, the sender would trigger the fast
   retransmit algorithm in the TCP sender, in response to the duplicate
   ACKs.  Research presented in [MV98] proposes the idea of suppressing
   or delaying the duplicate ACKs in the reverse direction to counteract
   this behavior.  Alternatively, proposals that make TCP more robust in
   the face of re-ordered segment arrivals [Flo99] may reduce the side
   effects of the re-ordering caused by link-layer retransmissions.

   A more high-level approach, outlined in the [DMT96], uses a new
   "corruption experienced" ICMP error message generated by routers that
   detect corruption.  These messages are sent in the forward direction,
   toward the packet's destination, rather than in the reverse direction
   as is done with ICMP Source Quench messages.  Sending the error
   messages in the forward direction allows this feedback to work over
   asymmetric paths.  As noted above, generating an error message in
   response to a damaged packet is problematic because the source and
   destination addresses may not be valid.  The mechanism outlined in
   [DMT96] gets around this problem by having the routers maintain a
   small cache of recent packet destinations; when the router
   experiences an error rate above some threshold, it sends an ICMP
   corruption-experienced message to all of the destinations in its
   cache.  Each TCP receiver then must return this information to its
   respective TCP sender (through a TCP option).  Upon receiving an ACK
   with this "corruption-experienced" option, the TCP sender assumes
   that packet loss is due to corruption rather than congestion for two
   round trip times (RTT) or until it receives additional link state
   information (such as "link down", source quench, or additional
   "corruption experienced" messages).  Note that in shared networks,
   ignoring segment loss for 2 RTTs may aggravate congestion by making
   TCP unresponsive.

3.3.4.3 Implementation Issues

   All of the techniques discussed above require changes to at least the
   TCP sending and receiving stacks, as well as intermediate routers.
   Due to the concerns over possibly ignoring congestion signals (i.e.,
   segment drops), the above algorithm is not recommended for use in
   shared networks.

3.3.4.4 Topology Considerations

   It is expected that corruption detection, in general would be
   beneficial in all environments outlined in section 2.  It would be
   particularly beneficial in the satellite/wireless environment over
   which these errors may be more prevalent.



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3.3.4.5 Possible Interaction and Relationships with Other Research

   SACK-based loss recovery algorithms (as described in 3.3.2) may
   reduce the impact of corrupted segments on mostly clean links because
   recovery will be able to happen more rapidly (and without relying on
   the retransmission timer).  Note that while SACK-based loss recovery
   helps, throughput will still suffer in the face of non-congestion
   related packet loss.

3.4 Congestion Avoidance

3.4.1  Mitigation Description

   During congestion avoidance, in the absence of loss, the TCP sender
   adds approximately one segment to its congestion window during each
   RTT [Jac88,Ste97,APS99].  Several researchers have observed that this
   policy leads to unfair sharing of bandwidth when multiple connections
   with different RTTs traverse the same bottleneck link, with the long
   RTT connections obtaining only a small fraction of their fair share
   of the bandwidth.

   One effective solution to this problem is to deploy fair queueing and
   TCP-friendly buffer management in network routers [Sut98].  However,
   in the absence of help from the network, other researchers have
   investigated changes to the congestion avoidance policy at the TCP
   sender, as described in [Flo91,HK98].

3.4.2 Research

   The "Constant-Rate" increase policy has been studied in [Flo91,HK98].
   It attempts to equalize the rate at which TCP senders increase their
   sending rate during congestion avoidance.  Both [Flo91] and [HK98]
   illustrate cases in which the "Constant-Rate" policy largely corrects
   the bias against long RTT connections, although [HK98] presents some
   evidence that such a policy may be difficult to incrementally deploy
   in an operational network.  The proper selection of a constant (for
   the constant rate of increase) is an open issue.

   The "Increase-by-K" policy can be selectively used by long RTT
   connections in a heterogeneous environment.  This policy simply
   changes the slope of the linear increase, with connections over a
   given RTT threshold adding "K" segments to the congestion window
   every RTT, instead of one.  [HK98] presents evidence that this
   policy, when used with small values of "K", may be successful in
   reducing the unfairness while keeping the link utilization high, when
   a small number of connections share a bottleneck link.  The selection
   of the constant "K," the RTT threshold to invoke this policy, and
   performance under a large number of flows are all open issues.



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3.4.3 Implementation Issues

   Implementation of either the "Constant-Rate" or "Increase-by-K"
   policies requires a change to the congestion avoidance mechanism at
   the TCP sender.  In the case of "Constant-Rate," such a change must
   be implemented globally.  Additionally, the TCP sender must have a
   reasonably accurate estimate of the RTT of the connection.  The
   algorithms outlined above violate the congestion avoidance algorithm
   as outlined in RFC 2581 [APS99] and therefore should not be
   implemented in shared networks at this time.

3.4.4 Topology Considerations

   These solutions are applicable to all satellite networks that are
   integrated with a terrestrial network, in which satellite connections
   may be competing with terrestrial connections for the same bottleneck
   link.

3.4.5 Possible Interaction and Relationships with Other Research

   As shown in [PADHV99], increasing the congestion window by multiple
   segments per RTT can cause TCP to drop multiple segments and force a
   retransmission timeout in some versions of TCP.  Therefore, the above
   changes to the congestion avoidance algorithm may need to be
   accompanied by a SACK-based loss recovery algorithm that can quickly
   repair multiple dropped segments.

3.5 Multiple Data Connections

3.5.1 Mitigation Description

   One method that has been used to overcome TCP's inefficiencies in the
   satellite environment is to use multiple TCP flows to transfer a
   given file.  The use of N TCP connections makes the sender N times
   more aggressive and therefore can improve throughput in some
   situations.  Using N multiple TCP connections can impact the transfer
   and the network in a number of ways, which are listed below.

   1. The transfer is able to start transmission using an effective
      congestion window of N segments, rather than a single segment as
      one TCP flow uses.  This allows the transfer to more quickly
      increase the effective cwnd size to an appropriate size for the
      given network.  However, in some circumstances an initial window
      of N segments is inappropriate for the network conditions.  In
      this case, a transfer utilizing more than one connection may
      aggravate congestion.





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   2. During the congestion avoidance phase, the transfer increases the
      effective cwnd by N segments per RTT, rather than the one segment
      per RTT increase that a single TCP connection provides.  Again,
      this can aid the transfer by more rapidly increasing the effective
      cwnd to an appropriate point.  However, this rate of increase can
      also be too aggressive for the network conditions.  In this case,
      the use of multiple data connections can aggravate congestion in
      the network.

   3. Using multiple connections can provide a very large overall
      congestion window.  This can be an advantage for TCP
      implementations that do not support the TCP window scaling
      extension [JBB92].  However, the aggregate cwnd size across all N
      connections is equivalent to using a TCP implementation that
      supports large windows.

   4. The overall cwnd decrease in the face of dropped segments is
      reduced when using N parallel connections.  A single TCP
      connection reduces the effective size of cwnd to half when a
      single segment loss is detected.  When utilizing N connections
      each using a window of W bytes, a single drop reduces the window
      to:

        (N * W) - (W / 2)

   Clearly this is a less dramatic reduction in the effective cwnd size
   than when using a single TCP connection.  And, the amount by which
   the cwnd is decreased is further reduced by increasing N.

   The use of multiple data connections can increase the ability of
   non-SACK TCP implementations to quickly recover from multiple dropped
   segments without resorting to a timeout, assuming the dropped
   segments cross connections.

   The use of multiple parallel connections makes TCP overly aggressive
   for many environments and can contribute to congestive collapse in
   shared networks [FF99].  The advantages provided by using multiple
   TCP connections are now largely provided by TCP extensions (larger
   windows, SACKs, etc.).  Therefore, the use of a single TCP connection
   is more "network friendly" than using multiple parallel connections.
   However, using multiple parallel TCP connections may provide
   performance improvement in private networks.









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3.5.2 Research

   Research on the use of multiple parallel TCP connections shows
   improved performance [IL92,Hah94,AOK95,AKO96].  In addition, research
   has shown that multiple TCP connections can outperform a single
   modern TCP connection (with large windows and SACK) [AHKO97].
   However, these studies did not consider the impact of using multiple
   TCP connections on competing traffic.  [FF99] argues that using
   multiple simultaneous connections to transfer a given file may lead
   to congestive collapse in shared networks.

3.5.3 Implementation Issues

   To utilize multiple parallel TCP connections a client application and
   the corresponding server must be customized.  As outlined in [FF99]
   using multiple parallel TCP connections is not safe (from a
   congestion control perspective) in shared networks and should not be
   used.

3.5.4 Topological Considerations

   As stated above, [FF99] outlines that the use of multiple parallel
   connections in a shared network, such as the Internet, may lead to
   congestive collapse.  However, the use of multiple connections may be
   safe and beneficial in private networks.  The specific topology being
   used will dictate the number of parallel connections required.  Some
   work has been done to determine the appropriate number of connections
   on the fly [AKO96], but such a mechanism is far from complete.

3.5.5 Possible Interaction and Relationships with Other Research

   Using multiple concurrent TCP connections enables use of a large
   congestion window, much like the TCP window scaling option [JBB92].
   In addition, a larger initial congestion window is achieved, similar
   to using [AFP98] or TCB sharing (see section 3.8).

3.6 Pacing TCP Segments

3.6.1 Mitigation Description

   Slow-start takes several round trips to fully open the TCP congestion
   window over routes with high bandwidth-delay products.  For short TCP
   connections (such as WWW traffic with HTTP/1.0), the slow-start
   overhead can preclude effective use of the high-bandwidth satellite
   links.  When senders implement slow-start restart after a TCP
   connection goes idle (suggested by Jacobson and Karels [JK92]),





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   performance is reduced in long-lived (but bursty) connections (such
   as HTTP/1.1, which uses persistent TCP connections to transfer
   multiple WWW page elements) [Hei97a].

   Rate-based pacing (RBP) is a technique, used in the absence of
   incoming ACKs, where the data sender temporarily paces TCP segments
   at a given rate to restart the ACK clock.  Upon receipt of the first
   ACK, pacing is discontinued and normal TCP ACK clocking resumes.  The
   pacing rate may either be known from recent traffic estimates (when
   restarting an idle connection or from recent prior connections), or
   may be known through external means (perhaps in a point-to-point or
   point-to-multipoint satellite network where available bandwidth can
   be assumed to be large).

   In addition, pacing data during the first RTT of a transfer may allow
   TCP to make effective use of high bandwidth-delay links even for
   short transfers.  However, in order to pace segments during the first
   RTT a TCP will have to be using a non-standard initial congestion
   window and a new mechanism to pace outgoing segments rather than send
   them back-to-back.  Determining an appropriate size for the initial
   cwnd is an open research question.  Pacing can also be used to reduce
   bursts in general (due to buggy TCPs or byte counting, see section
   3.2.2 for a discussion on byte counting).

3.6.2 Research

   Simulation studies of rate-paced pacing for WWW-like traffic have
   shown reductions in router congestion and drop rates [VH97a].  In
   this environment, RBP substantially improves performance compared to
   slow-start-after-idle for intermittent senders, and it slightly
   improves performance over burst-full-cwnd-after-idle (because of
   drops) [VH98].  More recently, pacing has been suggested to eliminate
   burstiness in networks with ACK filtering [BPK97].

3.6.3 Implementation Issues

   RBP requires only sender-side changes to TCP.  Prototype
   implementations of RBP are available [VH97b].  RBP requires an
   additional sender timer for pacing.  The overhead of timer-driven
   data transfer is often considered too high for practical use.
   Preliminary experiments suggest that in RBP this overhead is minimal
   because RBP only requires this timer for one RTT of transmission
   [VH98].  RBP is expected to make TCP more conservative in sending
   bursts of data after an idle period in hosts that do not revert to
   slow start after an idle period.  On the other hand, RBP makes TCP
   more aggressive if the sender uses the slow start algorithm to start
   the ACK clock after a long idle period.




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3.6.4  Topology Considerations

   RBP could be used to restart idle TCP connections for all topologies
   in Section 2.  Use at the beginning of new connections would be
   restricted to topologies where available bandwidth can be estimated
   out-of-band.

3.6.5 Possible Interaction and Relationships with Other Research

   Pacing segments may benefit from sharing state amongst various flows
   between two hosts, due to the time required to determine the needed
   information.  Additionally, pacing segments, rather than sending
   back-to-back segments, may make estimating the available bandwidth
   (as outlined in section 3.2.4) more difficult.

3.7 TCP Header Compression

   The TCP and IP header information needed to reliably deliver packets
   to a remote site across the Internet can add significant overhead,
   especially for interactive applications.  Telnet packets, for
   example, typically carry only a few bytes of data per packet, and
   standard IPv4/TCP headers add at least 40 bytes to this; IPv6/TCP
   headers add at least 60 bytes.  Much of this information remains
   relatively constant over the course of a session and so can be
   replaced by a short session identifier.

3.7.1 Mitigation Description

   Many fields in the TCP and IP headers either remain constant during
   the course of a session, change very infrequently, or can be inferred
   from other sources.  For example, the source and destination
   addresses, as well as the IP version, protocol, and port fields
   generally do not change during a session.  Packet length can be
   deduced from the length field of the underlying link layer protocol
   provided that the link layer packet is not padded.  Packet sequence
   numbers in a forward data stream generally change with every packet,
   but increase in a predictable manner.

   The TCP/IP header compression methods described in
   [DNP99,DENP97,Jac90] reduce the overhead of TCP sessions by replacing
   the data in the TCP and IP headers that remains constant, changes
   slowly, or changes in a predictable manner with a short "connection
   number".  Using this method, the sender first sends a full TCP/IP
   header, including in it a connection number that the sender will use
   to reference the connection.  The receiver stores the full header and
   uses it as a template, filling in some fields from the limited





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   information contained in later, compressed headers.  This compression
   can reduce the size of an IPv4/TCP headers from 40 to as few as 3 to
   5 bytes (3 bytes for some common cases, 5 bytes in general).

   Compression and decompression generally happen below the IP layer, at
   the end-points of a given physical link (such as at two routers
   connected by a serial line).  The hosts on either side of the
   physical link must maintain some state about the TCP connections that
   are using the link.

   The decompresser must pass complete, uncompressed packets to the IP
   layer.  Thus header compression is transparent to routing, for
   example, since an incoming packet with compressed headers is expanded
   before being passed to the IP layer.

   A variety of methods can be used by the compressor/decompressor to
   negotiate the use of header compression.  For example, the PPP serial
   line protocol allows for an option exchange, during which time the
   compressor/decompressor agree on whether or not to use header
   compression.  For older SLIP implementations, [Jac90] describes a
   mechanism that uses the first bit in the IP packet as a flag.

   The reduction in overhead is especially useful when the link is
   bandwidth-limited such as terrestrial wireless and mobile satellite
   links, where the overhead associated with transmitting the header
   bits is nontrivial.  Header compression has the added advantage that
   for the case of uniformly distributed bit errors, compressing TCP/IP
   headers can provide a better quality of service by decreasing the
   packet error probability.  The shorter, compressed packets are less
   likely to be corrupted, and the reduction in errors increases the
   connection's throughput.

   Extra space is saved by encoding changes in fields that change
   relatively slowly by sending only their difference from their values
   in the previous packet instead of their absolute values.  In order to
   decode headers compressed this way, the receiver keeps a copy of each
   full, reconstructed TCP header after it is decoded, and applies the
   delta values from the next decoded compressed header to the
   reconstructed full header template.

   A disadvantage to using this delta encoding scheme where values are
   encoded as deltas from their values in the previous packet is that if
   a single compressed packet is lost, subsequent packets with
   compressed headers can become garbled if they contain fields which
   depend on the lost packet.  Consider a forward data stream of packets
   with compressed headers and increasing sequence numbers.  If packet N
   is lost, the full header of packet N+1 will be reconstructed at the
   receiver using packet N-1's full header as a template.  Thus the



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   sequence number, which should have been calculated from packet N's
   header, will be wrong, the checksum will fail, and the packet will be
   discarded.  When the sending TCP times out and retransmits a packet
   with a full header is forwarded to re-synchronize the decompresser.

   It is important to note that the compressor does not maintain any
   timers, nor does the decompresser know when an error occurred (only
   the receiving TCP knows this, when the TCP checksum fails).  A single
   bit error will cause the decompresser to lose sync, and subsequent
   packets with compressed headers will be dropped by the receiving TCP,
   since they will all fail the TCP checksum. When this happens, no
   duplicate acknowledgments will be generated, and the decompresser can
   only re-synchronize when it receives a packet with an uncompressed
   header.  This means that when header compression is being used, both
   fast retransmit and selective acknowledgments will not be able
   correct packets lost on a compressed link.  The "twice" algorithm,
   described below, may be a partial solution to this problem.

   [DNP99] and [DENP97] describe TCP/IPv4 and TCP/IPv6 compression
   algorithms including compressing the various IPv6 extension headers
   as well as methods for compressing non-TCP streams.  [DENP97] also
   augments TCP header compression by introducing the "twice" algorithm.
   If a particular packet fails to decompress properly, the twice
   algorithm modifies its assumptions about the inferred fields in the
   compressed header, assuming that a packet identical to the current
   one was dropped between the last correctly decoded packet and the
   current one.  Twice then tries to decompress the received packet
   under the new assumptions and, if the checksum passes, the packet is
   passed to IP and the decompresser state has been re-synchronized.
   This procedure can be extended to three or more decoding attempts.
   Additional robustness can be achieved by caching full copies of
   packets which don't decompress properly in the hopes that later
   arrivals will fix the problem.  Finally, the performance improvement
   if the decompresser can explicitly request a full header is
   discussed.  Simulation results show that twice, in conjunction with
   the full header request mechanism, can improve throughput over
   uncompressed streams.

3.7.2 Research

   [Jac90] outlines a simple header compression scheme for TCP/IP.

   In [DENP97] the authors present the results of simulations showing
   that header compression is advantageous for both low and medium
   bandwidth links.  Simulations show that the twice algorithm, combined
   with an explicit header request mechanism, improved throughput by
   10-15% over uncompressed sessions across a wide range of bit error
   rates.



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   Much of this improvement may have been due to the twice algorithm
   quickly re-synchronizing the decompresser when a packet is lost.
   This is because the twice algorithm, applied one or two times when
   the decompresser becomes unsynchronized, will re-sync the
   decompresser in between 83% and 99% of the cases examined.  This
   means that packets received correctly after twice has resynchronized
   the decompresser will cause duplicate acknowledgments.  This re-
   enables the use of both fast retransmit and SACK in conjunction with
   header compression.

3.7.3 Implementation Issues

   Implementing TCP/IP header compression requires changes at both the
   sending (compressor) and receiving (decompresser) ends of each link
   that uses compression.  The twice algorithm requires very little
   extra machinery over and above header compression, while the explicit
   header request mechanism of [DENP97] requires more extensive
   modifications to the sending and receiving ends of each link that
   employs header compression.  Header compression does not violate
   TCP's congestion control mechanisms and therefore can be safely
   implemented in shared networks.

3.7.4 Topology Considerations

   TCP/IP header compression is applicable to all of the environments
   discussed in section 2, but will provide relatively more improvement
   in situations where packet sizes are small (i.e., overhead is large)
   and there is medium to low bandwidth and/or higher BER. When TCP's
   congestion window size is large, implementing the explicit header
   request mechanism, the twice algorithm, and caching packets which
   fail to decompress properly becomes more critical.

3.7.5 Possible Interaction and Relationships with Other Research

   As discussed above, losing synchronization between a sender and
   receiver can cause many packet drops.  The frequency of losing
   synchronization and the effectiveness of the twice algorithm may
   point to using a SACK-based loss recovery algorithm to reduce the
   impact of multiple lost segments.  However, even very robust SACK-
   based algorithms may not work well if too many segments are lost.

3.8 Sharing TCP State Among Similar Connections









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3.8.1 Mitigation Description

   Persistent TCP state information can be used to overcome limitations
   in the configuration of the initial state, and to automatically tune
   TCP to environments using satellite links and to coordinate multiple
   TCP connections sharing a satellite link.

   TCP includes a variety of parameters, many of which are set to
   initial values which can severely affect the performance of TCP
   connections traversing satellite links, even though most TCP
   parameters are adjusted later after the connection is established.
   These parameters include initial size of cwnd and initial MSS size.
   Various suggestions have been made to change these initial
   conditions, to more effectively support satellite links.  However, it
   is difficult to select any single set of parameters which is
   effective for all environments.

   An alternative to attempting to select these parameters a-priori is
   sharing state across TCP connections and using this state when
   initializing a new connection.  For example, if all connections to a
   subnet result in extended congestion windows of 1 megabyte, it is
   probably more efficient to start new connections with this value,
   than to rediscover it by requiring the cwnd to increase using slow
   start over a period of dozens of round-trip times.

3.8.2 Research

   Sharing state among connections brings up a number of questions such
   as what information to share, with whom to share, how to share it,
   and how to age shared information.  First, what information is to be
   shared must be determined.  Some information may be appropriate to
   share among TCP connections, while some information sharing may be
   inappropriate or not useful.  Next, we need to determine with whom to
   share information.  Sharing may be appropriate for TCP connections
   sharing a common path to a given host.  Information may be shared
   among connections within a host, or even among connections between
   different hosts, such as hosts on the same LAN.  However, sharing
   information between connections not traversing the same network may
   not be appropriate.  Given the state to share and the parties that
   share it, a mechanism for the sharing is required.  Simple state,
   like MSS and RTT, is easy to share, but congestion window information
   can be shared a variety of ways. The sharing mechanism determines
   priorities among the sharing connections, and a variety of fairness
   criteria need to be considered.  Also, the mechanisms by which
   information is aged require further study.  See RFC 2140 for a
   discussion of the security issues in both sharing state within a
   single host and sharing state among hosts on a subnet.  Finally, the
   security concerns associated with sharing a piece of information need



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   to be carefully considered before introducing such a mechanism.  Many
   of these open research questions must be answered before state
   sharing can be widely deployed.

   The opportunity for such sharing, both among a sequence of
   connections, as well as among concurrent connections, is described in
   more detail in [Tou97].  The state management itself is largely an
   implementation issue, however what information should be shared and
   the specific ways in which the information should be shared is an
   open question.

   Sharing parts of the TCB state was originally documented in T/TCP
   [Bra92], and is used there to aggregate RTT values across connection
   instances, to provide meaningful average RTTs, even though most
   connections are expected to persist for only one RTT.  T/TCP also
   shares a connection identifier, a sequence number separate from the
   window number and address/port pairs by which TCP connections are
   typically distinguished. As a result of this shared state, T/TCP
   allows a receiver to pass data in the SYN segment to the receiving
   application, prior to the completion of the three-way handshake,
   without compromising the integrity of the connection. In effect, this
   shared state caches a partial handshake from the previous connection,
   which is a variant of the more general issue of TCB sharing.

   Sharing state among connections (including transfers using non-TCP
   protocols) is further investigated in [BRS99].

3.8.3 Implementation Issues

   Sharing TCP state across connections requires changes to the sender's
   TCP stack, and possibly the receiver's TCP stack (as in the case of
   T/TCP, for example).  Sharing TCP state may make a particular TCP
   connection more aggressive.  However, the aggregate traffic should be
   more conservative than a group of independent TCP connections.
   Therefore, sharing TCP state should be safe for use in shared
   networks.  Note that state sharing does not present any new security
   problems within multiuser hosts.  In such a situation, users can
   steal network resources from one another with or without state
   sharing.

3.8.4 Topology Considerations

   It is expected that sharing state across TCP connections may be
   useful in all network environments presented in section 2.







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3.8.5 Possible Interaction and Relationships with Other Research

   The state sharing outlined above is very similar to the Congestion
   Manager proposal [BRS99] that attempts to share congestion control
   information among both TCP and UDP flows between a pair of hosts.

3.9 ACK Congestion Control

   In highly asymmetric networks, a low-speed return link can restrict
   the performance of the data flow on a high-speed forward link by
   limiting the flow of acknowledgments returned to the data sender.
   For example, if the data sender uses 1500 byte segments, and the
   receiver generates 40 byte acknowledgments (IPv4, TCP without
   options), the reverse link will congest with ACKs for asymmetries of
   more than 75:1 if delayed ACKs are used, and 37:1 if every segment is
   acknowledged.  For a 1.5 Mb/second data link, ACK congestion will
   occur for reverse link speeds below 20 kilobits/sec.  These levels of
   asymmetry will readily occur if the reverse link is shared among
   multiple satellite receivers, as is common in many VSAT satellite
   networks.  If a terrestrial modem link is used as a reverse link, ACK
   congestion is also likely, especially as the speed of the forward
   link is increased.  Current congestion control mechanisms are aimed
   at controlling the flow of data segments, but do not affect the flow
   of ACKs.

   In [KVR98] the authors point out that the flow of acknowledgments can
   be restricted on the low-speed link not only by the bandwidth of the
   link, but also by the queue length of the router.  The router may
   limit its queue length by counting packets, not bytes, and therefore
   begin discarding ACKs even if there is enough bandwidth to forward
   them.

3.9.1 Mitigation Description

   ACK Congestion Control extends the concept of flow control for data
   segments to acknowledgment segments.  In the method described in
   [BPK97], any intermediate router can mark an acknowledgment with an
   Explicit Congestion Notification (ECN) bit once the queue occupancy
   in the router exceeds a given threshold.  The data sender (which
   receives the acknowledgment) must "echo" the ECN bit back to the data
   receiver (see section 3.3.3 for a more detailed discussion of ECN).
   The proposed algorithm for marking ACK segments with an ECN bit is
   Random Early Detection (RED) [FJ93].  In response to the receipt of
   ECN marked data segments, the receiver will dynamically reduce the
   rate of acknowledgments using a multiplicative backoff.  Once
   segments without ECN are received, the data receiver speeds up
   acknowledgments using a linear increase, up to a rate of either 1 (no




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   delayed ACKs) or 2 (normal delayed ACKs) data segments per ACK.  The
   authors suggest that an ACK be generated at least once per window,
   and ideally a few times per window.

   As in the RED congestion control mechanism for data flow, the
   bottleneck gateway can randomly discard acknowledgments, rather than
   marking them with an ECN bit, once the queue fills beyond a given
   threshold.

3.9.2 Research

   [BPK97] analyze the effect of ACK Congestion Control (ACC) on the
   performance of an asymmetric network.  They note that the use of ACC,
   and indeed the use of any scheme which reduces the frequency of
   acknowledgments, has potential unwanted side effects.  Since each ACK
   will acknowledge more than the usual one or two data segments, the
   likelihood of segment bursts from the data sender is increased.  In
   addition, congestion window growth may be impeded if the receiver
   grows the window by counting received ACKs, as mandated by
   [Ste97,APS99].  The authors therefore combine ACC with a series of
   modifications to the data sender, referred to as TCP Sender
   Adaptation (SA).  SA combines a limit on the number of segments sent
   in a burst, regardless of window size.  In addition, byte counting
   (as opposed to ACK counting) is employed for window growth.  Note
   that byte counting has been studied elsewhere and can introduce
   side-effects, as well [All98].

   The results presented in [BPK97] indicate that using ACC and SA will
   reduce the bursts produced by ACK losses in unmodified (Reno) TCP.
   In cases where these bursts would lead to data loss at an
   intermediate router, the ACC and SA modification significantly
   improve the throughput for a single data transfer.  The results
   further suggest that the use of ACC and SA significantly improve
   fairness between two simultaneous transfers.

   ACC is further reported to prevent the increase in round trip time
   (RTT) that occurs when an unmodified TCP fills the reverse router
   queue with acknowledgments.

   In networks where the forward direction is expected to suffer losses
   in one of the gateways, due to queue limitations, the authors report
   at best a very slight improvement in performance for ACC and SA,
   compared to unmodified Reno TCP.








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3.9.3 Implementation Issues

   Both ACC and SA require modification of the sending and receiving
   hosts, as well as the bottleneck gateway.  The current research
   suggests that implementing ACC without the SA modifications results
   in a data sender which generates potentially disruptive segment
   bursts.  It should be noted that ACC does require host modifications
   if it is implemented in the way proposed in [BPK97].  The authors
   note that ACC can be implemented by discarding ACKs (which requires
   only a gateway modification, but no changes in the hosts), as opposed
   to marking them with ECN.  Such an implementation may, however,
   produce bursty data senders if it is not combined with a burst
   mitigation technique.  ACC requires changes to the standard ACKing
   behavior of a receiving TCP and therefore is not recommended for use
   in shared networks.

3.9.4 Topology Considerations

   Neither ACC nor SA require the storage of state in the gateway.
   These schemes should therefore be applicable for all topologies,
   provided that the hosts using the satellite or hybrid network can be
   modified.  However, these changes are expected to be especially
   beneficial to networks containing asymmetric satellite links.

3.9.5 Possible Interaction and Relationships with Other Research

   Note that ECN is a pre-condition for using ACK congestion control.
   Additionally, the ACK Filtering algorithm discussed in the next
   section attempts to solve the same problem as ACC.  Choosing between
   the two algorithms (or another mechanism) is currently an open
   research question.

3.10 ACK Filtering

   ACK Filtering (AF) is designed to address the same ACK congestion
   effects described in 3.9.  Contrary to ACC, however, AF is designed
   to operate without host modifications.

3.10.1 Mitigation Description

   AF takes advantage of the cumulative acknowledgment structure of TCP.
   The bottleneck router in the reverse direction (the low speed link)
   must be modified to implement AF.  Upon receipt of a segment which
   represents a TCP acknowledgment, the router scans the queue for
   redundant ACKs for the same connection, i.e. ACKs which acknowledge
   portions of the window which are included in the most recent ACK.
   All of these "earlier" ACKs are removed from the queue and discarded.




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   The router does not store state information, but does need to
   implement the additional processing required to find and remove
   segments from the queue upon receipt of an ACK.

3.10.2  Research

   [BPK97] analyzes the effects of AF.  As is the case in ACC, the use
   of ACK filtering alone would produce significant sender bursts, since
   the ACKs will be acknowledging more previously-unacknowledged data.
   The SA modifications described in 3.9.2 could be used to prevent
   those bursts, at the cost of requiring host modifications.  To
   prevent the need for modifications in the TCP stack, AF is more
   likely to be paired with the ACK Reconstruction (AR) technique, which
   can be implemented at the router where segments exit the slow reverse
   link.

   AR inspects ACKs exiting the link, and if it detects large "gaps" in
   the ACK sequence, it generates additional ACKs to reconstruct an
   acknowledgment flow which more closely resembles what the data sender
   would have seen had ACK Filtering not been introduced.  AR requires
   two parameters; one parameter is the desired ACK frequency, while the
   second controls the spacing, in time, between the release of
   consecutive reconstructed ACKs.

   In [BPK97], the authors show the combination of AF and AR to increase
   throughput, in the networks studied, over both unmodified TCP and the
   ACC/SA modifications.  Their results also strongly suggest that the
   use of AF alone, in networks where congestion losses are expected,
   decreases performance (even below the level of unmodified TCP Reno)
   due to sender bursting.

   AF delays acknowledgments from arriving at the receiver by dropping
   earlier ACKs in favor of later ACKs.  This process can cause a slight
   hiccup in the transmission of new data by the TCP sender.

3.10.3 Implementation Issues

   Both ACK Filtering and ACK Reconstruction require only router
   modification.  However, the implementation of AR requires some
   storage of state information in the exit router.  While AF does not
   require storage of state information, its use without AR (or SA)
   could produce undesired side effects.  Furthermore, more research is
   required regarding appropriate ranges for the parameters needed in
   AR.







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3.10.4 Topology Considerations

   AF and AR appear applicable to all topologies, assuming that the
   storage of state information in AR does not prove to be prohibitive
   for routers which handle large numbers of flows.  The fact that TCP
   stack modifications are not required for AF/AR makes this approach
   attractive for hybrid networks and networks with diverse types of
   hosts.  These modifications, however, are expected to be most
   beneficial in asymmetric network paths.

   On the other hand, the implementation of AF/AR requires the routers
   to examine the TCP header, which prohibits their use in secure
   networks where IPSEC is deployed.  In such networks, AF/AR can be
   effective only inside the security perimeter of a private, or virtual
   private network, or in private networks where the satellite link is
   protected only by link-layer encryption (as opposed to IPSEC).  ACK
   Filtering is safe to use in shared networks (from a congestion
   control point-of-view), as the number of ACKs can only be reduced,
   which makes TCP less aggressive.  However, note that while TCP is
   less aggressive, the delays that AF induces (outlined above) can lead
   to larger bursts than would otherwise occur.

3.10.5 Possible Interaction and Relationships with Other Research

   ACK Filtering attempts to solve the same problem as ACK Congestion
   Control (as outlined in section 3.9).  Which of the two algorithms is
   more appropriate is currently an open research question.

4   Conclusions

   This document outlines TCP items that may be able to mitigate the
   performance problems associated with using TCP in networks containing
   satellite links.  These mitigations are not IETF standards track
   mechanisms and require more study before being recommended by the
   IETF.  The research community is encouraged to examine the above
   mitigations in an effort to determine which are safe for use in
   shared networks such as the Internet.

5   Security Considerations

   Several of the above sections noted specific security concerns which
   a given mitigation aggravates.

   Additionally, any form of wireless communication link is more
   susceptible to eavesdropping security attacks than standard wire-
   based links due to the relative ease with which an attacker can watch
   the network and the difficultly in finding attackers monitoring the
   network.



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6   Acknowledgments

   Our thanks to Aaron Falk and Sally Floyd, who provided very helpful
   comments on drafts of this document.

7   References

   [AFP98]   Allman, M., Floyd, S. and C. Partridge, "Increasing TCP's
             Initial Window", RFC 2414, September 1998.

   [AGS99]   Allman, M., Glover, D. and L. Sanchez, "Enhancing TCP Over
             Satellite Channels using Standard Mechanisms", BCP 28, RFC
             2488, January 1999.

   [AHKO97]  Mark Allman, Chris Hayes, Hans Kruse, Shawn Ostermann.  TCP
             Performance Over Satellite Links.  In Proceedings of the
             5th International Conference on Telecommunication Systems,
             March 1997.

   [AHO98]   Mark Allman, Chris Hayes, Shawn Ostermann.  An Evaluation
             of TCP with Larger Initial Windows.  Computer Communication
             Review, 28(3), July 1998.

   [AKO96]   Mark Allman, Hans Kruse, Shawn Ostermann.  An Application-
             Level Solution to TCP's Satellite Inefficiencies.  In
             Proceedings of the First International Workshop on
             Satellite-based Information Services (WOSBIS), November
             1996.

   [All97a]  Mark Allman.  Improving TCP Performance Over Satellite
             Channels.  Master's thesis, Ohio University, June 1997.

   [All97b]  Mark Allman.  Fixing Two BSD TCP Bugs.  Technical Report
             CR-204151, NASA Lewis Research Center, October 1997.

   [All98]   Mark Allman. On the Generation and Use of TCP
             Acknowledgments.  ACM Computer Communication Review, 28(5),
             October 1998.

   [AOK95]   Mark Allman, Shawn Ostermann, Hans Kruse.  Data Transfer
             Efficiency Over Satellite Circuits Using a Multi-Socket
             Extension to the File Transfer Protocol (FTP).  In
             Proceedings of the ACTS Results Conference, NASA Lewis
             Research Center, September 1995.

   [AP99]    Mark Allman, Vern Paxson.  On Estimating End-to-End Network
             Path Properties. ACM SIGCOMM, September 1999.




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   [APS99]   Allman, M., Paxson, V. and W. Richard Stevens, "TCP
             Congestion Control", RFC 2581, April 1999.

   [BCC+98]  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.

   [BKVP97]  B. Bakshi and P. Krishna and N. Vaidya and D. Pradham,
             "Improving Performance of TCP over Wireless Networks", 17th
             International Conference on Distributed Computing Systems
             (ICDCS), May 1997.

   [BPK97]   Hari Balakrishnan, Venkata N. Padmanabhan, and Randy H.
             Katz.  The Effects of Asymmetry on TCP Performance.  In
             Proceedings of the ACM/IEEE Mobicom, Budapest, Hungary,
             ACM.  September, 1997.

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

   [BPSK96]  H. Balakrishnan and V. Padmanabhan and S. Sechan and R.
             Katz, "A Comparison of Mechanisms for Improving TCP
             Performance over Wireless Links", ACM SIGCOMM, August 1996.

   [Bra89]   Braden, R., "Requirements for Internet Hosts --
             Communication Layers", STD 3, RFC 1122, October 1989.

   [Bra92]   Braden, R., "Transaction TCP -- Concepts", RFC 1379,
             September 1992.

   [Bra94]   Braden, R., "T/TCP -- TCP Extensions for Transactions:
             Functional Specification", RFC 1644, July 1994.

   [BRS99]   Hari Balakrishnan, Hariharan Rahul, and Srinivasan Seshan.
             An Integrated Congestion Management Architecture for
             Internet Hosts.  ACM SIGCOMM, September 1999.

   [ddKI99]  M. deVivo, G.O. deVivo, R. Koeneke, G. Isern.  Internet
             Vulnerabilities Related to TCP/IP and T/TCP.  Computer
             Communication Review, 29(1), January 1999.







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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


   [DENP97]  Mikael Degermark, Mathias Engan, Bjorn Nordgren, Stephen
             Pink.  Low-Loss TCP/IP Header Compression for Wireless
             Networks.  ACM/Baltzer Journal on Wireless Networks, vol.3,
             no.5, p. 375-87.

   [DMT96]   R. C. Durst and G. J. Miller and E. J. Travis, "TCP
             Extensions for Space Communications", Mobicom 96, ACM, USA,
             1996.

   [DNP99]   Degermark, M., Nordgren, B. and S. Pink, "IP Header
             Compression", RFC 2507, February 1999.

   [FF96]    Kevin Fall, Sally Floyd.  Simulation-based Comparisons of
             Tahoe, Reno, and SACK TCP.  Computer Communication Review,
             V. 26 N. 3, July 1996, pp. 5-21.

   [FF99]    Sally Floyd, Kevin Fall.  Promoting the Use of End-to-End
             Congestion Control in the Internet, IEEE/ACM Transactions
             on Networking, August 1999.

   [FH99]    Floyd, S. and T. Henderson, "The NewReno Modification to
             TCP's Fast Recovery Algorithm", RFC 2582, April 1999.

   [FJ93]    Sally Floyd and Van Jacobson.  Random Early Detection
             Gateways for Congestion Avoidance, IEEE/ACM Transactions on
             Networking, V. 1 N. 4, August 1993.

   [Flo91]   Sally Floyd.  Connections with Multiple Congested Gateways
             in Packet-Switched Networks, Part 1: One-way Traffic.  ACM
             Computer Communications Review, V. 21, N. 5, October 1991.

   [Flo94]   Sally Floyd.  TCP and Explicit Congestion Notification, ACM
             Computer Communication Review, V. 24 N. 5, October 1994.

   [Flo99]   Sally Floyd.  "Re: TCP and out-of-order delivery", email to
             end2end-interest mailing list, February, 1999.

   [Hah94]   Jonathan Hahn.  MFTP: Recent Enhancements and Performance
             Measurements.  Technical Report RND-94-006, NASA Ames
             Research Center, June 1994.

   [Hay97]   Chris Hayes.  Analyzing the Performance of New TCP
             Extensions Over Satellite Links.  Master's Thesis, Ohio
             University, August 1997.

   [HK98]    Tom Henderson, Randy Katz.  On Improving the Fairness of
             TCP Congestion Avoidance.  Proceedings of IEEE Globecom `98
             Conference, 1998.



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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


   [HK99]    Tim Henderson, Randy Katz.  Transport Protocols for
             Internet-Compatible Satellite Networks, IEEE Journal on
             Selected Areas of Communications, February, 1999.

   [Hoe95]   J. Hoe, Startup Dynamics of TCP's Congestion Control and
             Avoidance Schemes. Master's Thesis, MIT, 1995.

   [Hoe96]   Janey Hoe.  Improving the Startup Behavior of a Congestion
             Control Scheme for TCP.  In ACM SIGCOMM, August 1996.

   [IL92]    David Iannucci and John Lakashman.  MFTP: Virtual TCP
             Window Scaling Using Multiple Connections.  Technical
             Report RND-92-002, NASA Ames Research Center, January 1992.

   [Jac88]   Van Jacobson.  Congestion Avoidance and Control.  In
             Proceedings of the SIGCOMM '88, ACM.  August, 1988.

   [Jac90]   Jacobson, V., "Compressing TCP/IP Headers", RFC 1144,
             February 1990.

   [JBB92]   Jacobson, V., Braden, R. and D. Borman, "TCP Extensions for
             High Performance", RFC 1323, May 1992.

   [JK92]    Van Jacobson and Mike Karels.  Congestion Avoidance and
             Control.  Originally appearing in the proceedings of
             SIGCOMM '88 by Jacobson only, this revised version includes
             an additional appendix.  The revised version is available
             at ftp://ftp.ee.lbl.gov/papers/congavoid.ps.Z.  1992.

   [Joh95]   Stacy Johnson.  Increasing TCP Throughput by Using an
             Extended Acknowledgment Interval.  Master's Thesis, Ohio
             University, June 1995.

   [KAGT98]  Hans Kruse, Mark Allman, Jim Griner, Diepchi Tran.  HTTP
             Page Transfer Rates Over Geo-Stationary Satellite Links.
             March 1998. Proceedings of the Sixth International
             Conference on Telecommunication Systems.

   [Kes91]   Srinivasan Keshav.  A Control Theoretic Approach to Flow
             Control.  In ACM SIGCOMM, September 1991.

   [KM97]    S. Keshav, S. Morgan. SMART Retransmission: Performance
             with Overload and Random Losses. Proceeding of Infocom.
             1997.







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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


   [KVR98]   Lampros Kalampoukas, Anujan Varma, and K. K.Ramakrishnan.
             Improving TCP Throughput Over Two-Way Asymmetric Links:
             Analysis and Solutions.  Measurement and Modeling of
             Computer Systems, 1998, Pages 78-89.

   [MM96a]   M. Mathis, J. Mahdavi, "Forward Acknowledgment: Refining
             TCP Congestion Control," Proceedings of SIGCOMM'96, August,
             1996, Stanford, CA.  Available from
             http://www.psc.edu/networking/papers/papers.html

   [MM96b]   M. Mathis, J. Mahdavi, "TCP Rate-Halving with Bounding
             Parameters" Available from
             http://www.psc.edu/networking/papers/FACKnotes/current.

   [MMFR96]  Mathis, M., Mahdavi, J., Floyd, S. and A. Romanow, "TCP
             Selective Acknowledgment Options", RFC 2018, October 1996.

   [MSMO97]  M. Mathis, J. Semke, J. Mahdavi, T. Ott, "The Macroscopic
             Behavior of the TCP Congestion Avoidance
             Algorithm",Computer Communication Review, volume 27,
             number3, July 1997.  Available from
             http://www.psc.edu/networking/papers/papers.html

   [MV98]    Miten N. Mehta and Nitin H. Vaidya.  Delayed Duplicate-
             Acknowledgments: A Proposal to Improve Performance of TCP
             on Wireless Links.  Technical Report 98-006, Department of
             Computer Science, Texas A&M University, February 1998.

   [Nic97]   Kathleen Nichols.  Improving Network Simulation with
             Feedback.  Com21, Inc. Technical Report.  Available from
             http://www.com21.com/pages/papers/068.pdf.

   [PADHV99] Paxson, V., Allman, M., Dawson, S., Heavens, I. and B.
             Volz, "Known TCP Implementation Problems", RFC 2525, March
             1999.

   [Pax97]   Vern Paxson.  Automated Packet Trace Analysis of TCP
             Implementations.  In Proceedings of ACM SIGCOMM, September
             1997.

   [PN98]    Poduri, K. and K. Nichols, "Simulation Studies of Increased
             Initial TCP Window Size", RFC 2415, September 1998.

   [Pos81]   Postel, J., "Transmission Control Protocol", STD 7, RFC
             793, September 1981.






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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


   [RF99]    Ramakrishnan, K. and S. Floyd, "A Proposal to add Explicit
             Congestion Notification (ECN) to IP", RFC 2481, January
             1999.

   [SF98]    Nihal K. G. Samaraweera and Godred Fairhurst,
             "Reinforcement of TCP error Recovery for Wireless
             Communication", Computer Communication Review, volume 28,
             number 2, April 1998.

   [SP98]    Shepard, T. and C. Partridge, "When TCP Starts Up With Four
             Packets Into Only Three Buffers", RFC 2416, September 1998.

   [Ste97]   Stevens, W., "TCP Slow Start, Congestion Avoidance, Fast
             Retransmit, and Fast Recovery Algorithms", RFC 2001,
             January 1997.

   [Sut98]   B. Suter, T. Lakshman, D. Stiliadis, and A. Choudhury.
             Design Considerations for Supporting TCP with Per-flow
             Queueing.  Proceedings of IEEE Infocom `98 Conference,
             1998.

   [Tou97]   Touch, J., "TCP Control Block Interdependence", RFC 2140,
             April 1997.

   [VH97a]   Vikram Visweswaraiah and John Heidemann.  Improving Restart
             of Idle TCP Connections.  Technical Report 97-661,
             University of Southern California, 1997.

   [VH97b]   Vikram Visweswaraiah and John Heidemann.  Rate-based pacing
             Source Code Distribution, Web page:
             http://www.isi.edu/lsam/publications/rate_based_pacing/README.html
             November, 1997.

   [VH98]    Vikram Visweswaraiah and John Heidemann.  Improving Restart
             of Idle TCP Connections (revised).  Submitted for
             publication.















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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


8   Authors' Addresses

   Mark Allman
   NASA Glenn Research Center/BBN Technologies
   Lewis Field
   21000 Brookpark Rd.  MS 54-2
   Cleveland, OH  44135

   EMail: mallman@grc.nasa.gov
   http://roland.grc.nasa.gov/~mallman


   Spencer Dawkins
   Nortel
   P.O.Box 833805
   Richardson, TX 75083-3805

   EMail: Spencer.Dawkins.sdawkins@nt.com


   Dan Glover
   NASA Glenn Research Center
   Lewis Field
   21000 Brookpark Rd.  MS 3-6
   Cleveland, OH  44135

   EMail: Daniel.R.Glover@grc.nasa.gov
   http://roland.grc.nasa.gov/~dglover


   Jim Griner
   NASA Glenn Research Center
   Lewis Field
   21000 Brookpark Rd.  MS 54-2
   Cleveland, OH  44135

   EMail: jgriner@grc.nasa.gov
   http://roland.grc.nasa.gov/~jgriner


   Diepchi Tran
   NASA Glenn Research Center
   Lewis Field
   21000 Brookpark Rd.  MS 54-2
   Cleveland, OH  44135

   EMail: dtran@grc.nasa.gov




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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


   Tom Henderson
   University of California at Berkeley
   Phone: +1 (510) 642-8919

   EMail: tomh@cs.berkeley.edu
   URL: http://www.cs.berkeley.edu/~tomh/


   John Heidemann
   University of Southern California/Information Sciences Institute
   4676 Admiralty Way
   Marina del Rey, CA 90292-6695

   EMail: johnh@isi.edu


   Joe Touch
   University of Southern California/Information Sciences Institute
   4676 Admiralty Way
   Marina del Rey, CA 90292-6601
   USA

   Phone: +1 310-448-9151
   Fax:   +1 310-823-6714
   URL:   http://www.isi.edu/touch
   EMail: touch@isi.edu


   Hans Kruse
   J. Warren McClure School of Communication Systems Management
   Ohio University
   9 S. College Street
   Athens, OH 45701

   Phone: 740-593-4891
   Fax: 740-593-4889
   EMail: hkruse1@ohiou.edu
   http://www.csm.ohiou.edu/kruse


   Shawn Ostermann
   School of Electrical Engineering and Computer Science
   Ohio University
   416 Morton Hall
   Athens, OH  45701

   Phone: (740) 593-1234
   EMail: ostermann@cs.ohiou.edu



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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


   Keith Scott
   The MITRE Corporation
   M/S W650
   1820 Dolley Madison Blvd.
   McLean VA 22102-3481

   EMail: kscott@mitre.org


   Jeffrey Semke
   Pittsburgh Supercomputing Center
   4400 Fifth Ave.
   Pittsburgh, PA  15213

   EMail: semke@psc.edu
   http://www.psc.edu/~semke



































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RFC 2760       Ongoing TCP Research Related to Satellites  February 2000


9  Full Copyright Statement

   Copyright (C) The Internet Society (2000).  All Rights Reserved.

   This document and translations of it may be copied and furnished to
   others, and derivative works that comment on or otherwise explain it
   or assist in its implementation may be prepared, copied, published
   and distributed, in whole or in part, without restriction of any
   kind, provided that the above copyright notice and this paragraph are
   included on all such copies and derivative works.  However, this
   document itself may not be modified in any way, such as by removing
   the copyright notice or references to the Internet Society or other
   Internet organizations, except as needed for the purpose of
   developing Internet standards in which case the procedures for
   copyrights defined in the Internet Standards process must be
   followed, or as required to translate it into languages other than
   English.

   The limited permissions granted above are perpetual and will not be
   revoked by the Internet Society or its successors or assigns.

   This document and the information contained herein is provided on an
   "AS IS" basis and THE INTERNET SOCIETY AND THE INTERNET ENGINEERING
   TASK FORCE DISCLAIMS 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.

Acknowledgement

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



















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