One document matched: draft-ietf-tcpsat-res-issues-04.txt
Differences from draft-ietf-tcpsat-res-issues-03.txt
Internet Engineering Task Force Mark Allman, Editor
INTERNET DRAFT Spencer Dawkins
File: draft-ietf-tcpsat-res-issues-04.txt Dan Glover
Jim Griner
John Heidemann
Shawn Ostermann
Keith Scott
Jeffrey Semke
Joe Touch
Diepchi Tran
August, 1998
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Ongoing TCP Research Related to Satellites
Status of this Memo
This document is an Internet-Draft. Internet-Drafts are working
documents of the Internet Engineering Task Force (IETF), its areas,
and its working groups. Note that other groups may also distribute
working documents as Internet-Drafts.
Internet-Drafts are draft documents valid for a maximum of six
months and may be updated, replaced, or obsoleted by other documents
at any time. It is inappropriate to use Internet-Drafts as
reference material or to cite them other than as ``work in
progress.''
To learn the current status of any Internet-Draft, please check the
``1id-abstracts.txt'' listing contained in the Internet- Drafts
Shadow Directories on ftp.is.co.za (Africa), nic.nordu.net (Europe),
munnari.oz.au (Pacific Rim), ds.internic.net (US East Coast), or
ftp.isi.edu (US West Coast).
NOTE
This document is not to be taken as a finished product. Some of the
sections are rough and are included in order to obtain comments from
the community that will benefit future iterations of this document.
This is simply a step in the ongoing conversation about this
document. Finally, all the authors of this draft do not necessarily
agree with and/or advocate all the mechanisms outlined in this
document.
Abstract
This document outlines TCP mechanisms that may help better utilize
the available bandwidth in TCP transfers over long-delay satellite
channels. The work outlined in this document is preliminary and has
not yet been judged to be safe for use in the shared Internet. In
addition, some of the work outlined in this document has been shown
to be unsafe for shared networks, but may be acceptable for use in
private networks.
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Table of Contents
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. The proposals outlined in this document
are currently being studied throughout the research community.
Therefore, these mechanisms SHOULD NOT be used in the shared
Internet. If, at some point, the mechanisms discussed in this memo
prove safe and appropriate for general use, the appropriate IETF
documents will be written. Until that time, these mechanisms should
be used for research and in private networks only.
It should be noted that non-TCP mechanisms that help performance
over satellite channels do exist (e.g., application-level changes,
queueing disciplines, etc.). However, outlining these non-TCP
mitigations is left as future work.
2 Satellite Architectures
Satellite characteristics are discussed in [AG98]. This section
discusses several ways that satellites might be used in the
Internet.
2.1 Asymmetric Satellite Networks
Some satellite networks exhibit a bandwidth asymmetry, with a larger
data rate in one direction than the other, because of limits on the
transmission power and the antenna size at one end of the link.
Meanwhile, other satellite systems are one way only 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 channels.
2.2 Satellite Link as Last Hop
Satellite links that provide service directly to end users may allow
for specialized design of protocols used over the last hop. Some
satellite providers use the satellite channel as a shared high speed
downlink to users with a lower speed, non-shared terrestrial channel
that is used as a return channel for requests and acknowledgments.
Many times this creates an asymmetric network, as discussed in
section 2.1.
2.3 Hybrid Satellite Networks
In the more general case, satellites may be located at any point in
the network topology. In this case, the satellite link acts as just
another channel between two gateways. In this environment, a given
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connection may be sent over terrestrial channels (including
wireless), as well as satellite channels. 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 channel. There is no terrestrial traffic to
contend with in this environment. This pure satellite environment
exhibits only the problems associated with the satellite channels,
as outlined in [AG98]. Since this is a private network, some
mitigations to TCP's inefficiencies can be used that are not
suitable for shared networks, such as the Internet.
2.5 Point-to-Multipoint Satellite Networks
Satellites have a natural advantage in point-to-multipoint
communication. Although satellite communications began as a
trunking method for telephony, the broadcast advantages of
satellites were quickly recognized and utilized for television
program distribution. One signal can be transmitted up to a
satellite and then relayed back down to a large geographic area.
Any ground station in that area can pick up the signal if tuned to
the right channel. In the same way, data can be transmitted to
small ground stations located over large geographic distances
without loading terrestrial networks. Satellites have found use in
corporate intranets and VSAT (very small aperture terminal) networks
especially for database applications, but advantages for WWW
caching, distributing network news, and multicasting applications
are obvious and could help to reduce network congestion. While this
is a valuable use of satellite systems, it is considered out of
scope in this document, as TCP is a unicast-only protocol.
2.6 Multiple Satellite Hops
In some cases, service may be provided over multiple satellite hops.
This aggravates the satellite characteristics described in [AG98].
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. The
implementation issues of the mechanism will be discussed next.
Finally, the mechanism's benefits in each of the environments above
will be outlined.
3.1 TCP For Transactions
3.1.1 Mitigation Description
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TCP uses a three-way handshake to setup a connection between two
hosts [Pos81]. This connection setup requires 1-1.5 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]. In many situations,
T/TCP is able to bypass the three-way handshake. This allows 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] and [Bra94].
3.1.3 Implementation Issues
T/TCP requires changes in the TCP stacks of both the data sender and
the data receiver. There are some security implications of sending
data in the first data segment. These will be briefly presented
and/or pointed at in a future iteration of this document. In
addition, some researchers feel that the costs associated with
implementing T/TCP outweigh the potential benefits.
3.1.4 Topology Considerations
It is expected that T/TCP will be equally beneficial in all
environments outlined in section 2.
3.2 Slow Start
The slow start algorithm is used to gradually increase the size of
TCP's sliding window [Jac88] [Ste97]. 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 capacity [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] allows data receivers to refrain from
ACKing every incoming data segment. However, every second
full-sized segment must 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 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
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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 make the
initial value of cwnd be more than a single segment, as required by
[Bra89] and [Ste97]. An experimental TCP extension outlined in
[AFP98] allows the initial size of cwnd to be increased, according
to 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.
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
channels respectively. Furthermore, it has been shown that using an
initial cwnd of 4 packets does not negatively impact overall
performance over dialup modem channels with a small number of
buffers [SP97]. [AHO98] shows an improvement in transfer time for 16
KB files across the Internet and dialup modem channels when using a
larger initial value for cwnd. However, a slight increase in
retransmitted 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.
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.
3.2.2 Byte Counting
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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. One mechanism that can
mitigate this problems caused by delayed ACKs is the use of ``byte
counting'' [All97a] [All98]. Using this mechanism, the cwnd
increase is based on the number of previously unacknowledged bytes
ACKed, 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.
Byte counting leads to slightly larger line-rate bursts of segments.
This increase in burstiness may increase the loss rate on some
networks. The size of the line-rate burst increases if the receiver
generates ``stretch ACKs'' [Pax97] (either by design [Joh95] or due
to implementation bugs [All97b] [PADHV97]), since a stretch ACK
covers more previously unacknowledged bytes than a normal delayed
ACK. Therefore, the increase in cwnd when using byte counting may
cause an inappropriate line-rate burst. One way to prevent these
line-rate bursts is to use a ``limited byte counting'' mechanism, as
outlined in [All98]. In this form of byte counting, cwnd is
increased by the number of previously unacknowledged bytes ACKed by
each incoming ACK, but the increase is limited to 2 segments.
[All98] shows that this approach prevents large line-rate bursts
that hurt performance.
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 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.
3.2.2.3 Implementation Issues
Changing from ACK counting to byte counting requires changes to the
data sender's TCP stack.
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 mitigates the retarded window growth provided by
receivers that generate stretch ACKs because of the capacity of the
return channel, as discussed in [BPK97].
3.2.3 Disabling Delayed ACKs During Slow Start
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(in progress)
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 [Jac88] algorithm is called
"ssthresh". 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] 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 with 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
remains an open research area. Therefore, before this mechanism is
widely deployed, it must be studied in a more dynamic network
environment.
3.2.4.3 Implementation Issues
Estimating ssthresh requires changes to the data sender's TCP
stack.
3.2.4.4 Topology Considerations
It is expected that this mechanism will work well in all symmetric
topologies outlined in section 2. However, asymmetric channels 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.
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3.3 Loss Recovery
3.3.1 Non-SACK Based Mechanisms
(in progress)
3.3.2 SACK Based Mechanisms
3.3.2.1 SACK "pipe" Algorithm
(in progress)
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, in which one packet
is sent for every two ACKs received during recovery. ACKing
every-other packet 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].
In addition, the FACK with Rate-Halving implementation provides
Thresholded Retransmission to each lost segment. Tcprexmtthresh is
the number of duplicate ACKs required by Reno to 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.
3.3.2.2.2 Research
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The original FACK algorithm was presented at Sigcomm'96 [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 the theoretical maximum for TCP than either SACK or
Reno [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.
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 Reno, it may be considerably more attractive over long delay
paths.
3.3.3 Explicit Congestion Notification
3.3.3.1 Mitigation Description
Explicit congestion notification (ECN) allows routers to inform TCP
senders about congestion levels. Two major forms of ECN have been
studied. If a router employs backward ECN (BECN), it transmits
packets to the data originator informing them 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 cwnd value. The second major form of congestion
notification is forward ECN (FECN). In this form of ECN, routers
mark data segments when congestion is imminent, but forward the data
segment. The data receiver echos the congestion information back to
the sender in the ACK packet. A current IETF proposal specifies an
implementation of FECN [RF98].
As proposed in [RF98], TCP senders transmit segments with an
``ECN-capable'' bit set in the packet header. If a router employing
an active queueing strategy, such as Random Early Detection (RED)
[FJ93] [BCC+98], would otherwise drop this segment, an ``ECN
experienced'' bit is set instead. The TCP receiver echos this
information back to the sender in ACK segments. The TCP sender
reacts just as it would if a segment was dropped, and reduces its
congestion window appropriately.
A side-effect of implementing ECN, as suggested in [RF98], is that
the intervening routers will employ active queueing mechanisms.
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 router queue.
Since satellite networks generally have higher bit-error rates than
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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
complete solution to this problem, adding an ECN mechanism to TCP
may 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 in 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 FECN and BECN.
A proposal for implementing ECN in the Internet is currently being
discussed within the IETF [RF98].
3.3.3.3 Implementation Issues
Deployment of ECN requires changes to the TCP implementation on both
sender and receiver.
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 a ``marked'' segments while still notifying the end
nodes that congestion is occurring along the path.
3.3.3.4 Topology Considerations
It is expected that none of the environments outlined in section 2
will present a bias towards ECN traffic.
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], is to assume that all loss is due to congestion
and to trigger the congestion control algorithms, as defined in
[Ste97]. In the worst case, this loss is detected by the expiration
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of TCP's retransmission timer. Alternately, the loss might be
inferred from duplicate ACKs (by the fast retransmit algorithm
[Jac90] [Ste97]) or by a hole in a SACK block.
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 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 [AG98]. 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 terrestrial 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
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of view. As a simple example, if segments A B C D are sent across a
noisy channel and segment B is corrupted, segments C and D may have
already crossed the channel before B can be retransmitted at the
link 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, which
may trigger the fast retransmit algorithm in TCP sender, in some
cases. Research presented in [MV98] proposes the idea of
suppressing or delaying the DUPACKs in the reverse direction to
counteract this behavior.
A more high-level approach, outlined in the SCPS-TP work [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. SCPS 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 or until it receives
additional link state information (such as "link down", source
quench, or additional "corruption experienced" messages).
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.
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.
3.4 Spoofing
[Editor's Note: This section may be removed in the future, depending
upon the evolution of the "tcppep" initiative.]
3.4.1 Mitigation Description
TCP spoofing is a technique used to split a TCP connection between a
client (such as a mobile host or a hybrid terminal) and a server
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(such as fixed terminal or Internet server) into two parts: one
between the client and its gateway router over a satellite/wireless
link and the other between the gateway router and the server over
the Internet/wired link. The gateway effectively breaks incoming
TCP connections in two by acting on the client's behalf in
interactions with the server. This allows the server to complete
the transfer without incurring delays introduced by the satellite.
Furthermore, spoofing allows the gateway to use a more appropriate
transport protocol (or version of TCP) over the satellite hop. This
mechanism is criticized by some as breaking the end-to-end semantics
associated with the TCP protocol.
3.4.2 Research
The TCP spoofing technique has been used to improve the overall
throughput for asymmetric Internet access over satellite-terrestrial
network [ASBD96] and for transferring data to mobile clients over
wireless-wired network [BPSK97] [BB95]. In addition, [ASBD96] with
spoofing and an increased ACK interval (i.e., decreased frequency of
ACKs), it has been found that the throughput increased up to 400Kbps
compare to 120Kbps of the system without these techniques. By using
spoofing and the SMART retransmission technique [KM97], [BPSK97]
shows that the TCP throughput improved from 0.7 Mbps to 1.3 Mbps in
LAN environments and from 0.3 Mbps to 1.1 Mbps in WAN environments.
3.4.3 Implementation Issues
The use of TCP spoofing requires modification to the gateway
routers to enable them to act on the behalf of the end hosts.
3.4.4 Topology Considerations
TCP spoofing should help performance over all topologies outlined
above. However, TCP spoofing is an especially useful technique in
asymmetric networks.
3.5 snoop
[Editor's Note: This section may be removed in the future, depending
upon the evolution of the "tcppep" initiative.]
3.6 Multiple Data Connections
3.6.1 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 benefit 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
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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.
2. During the congestion avoidance phase, the transfer increases
the effective cwnd by N segments per RTT, rather than one
segment per RTT that a single TCP connection would. 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 cwnd
size. 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
segment loss is detected. Therefore, 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.
The use of multiple data connections can increase the ability of
non-SACK TCP implementations to quickly recover from multiple
dropped segments, 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 [FF98]. 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.
3.6.2 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
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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. [FF98] argues that using
multiple simultaneous connections to transfer a given file may lead
to congestive collapse in shared networks.
3.6.3 Implementation Issues
To utilize multiple parallel TCP connections a client application
and the corresponding server must be customized.
3.6.4 Topological Considerations
As stated above, [FF98] 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.
3.7 Pacing TCP Segments
3.7.1 ACK Spacing
3.7.1.1 Mitigation Description
Routes with high bandwidth*delay products (such as those found in
geostationary satellite links) are capable of utilizing large TCP
congestion windows. However, it can take a long time before TCP can
fully utilize this large window. One possible cause of this delay
are small router buffers, since in an idealized situation the router
buffer should be one half the bandwidth*delay product in order to
avoid losing segments [Par97]. This arises during slow start,
because it is possible for the sender to burst data at twice the
rate of the bottleneck router.
Using ACK spacing, the bursts can be spread over time by using a
separation of at least two segments between ACKs [Par97]. Since the
ACK rate is used to determine the rate packets are sent, ACK spacing
would allow the sender to transmit at the correct rate.
3.7.1.2 Research
Currently an implementation of ACK spacing does not exist, beyond a
mere thought exercise. An algorithm has not been developed to
determine the proper ACK spacing, which may be different depending
on whether TCP is in slow start or congestion avoidance.
3.7.1.3 Implementation Issues
ACK spacing can be implemented in the router, which elevates the
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need to change either the sender or receiver's TCP stack.
3.7.1.4 Topology Considerations
It may not be necessary to use ACK spacing in an asymmetrical
routes, because of their inherent nature.
3.7.2 Rate-Based Pacing
3.7.2.1 Mitigation Description
Slow-start takes several round trips to fully open the TCP
congestion window over routes with high bandwidth-delay product.
For short TCP connections (common in web traffic with HTTP/1.0),
this slow-start overhead can preclude effective use of the
high-bandwidth satellite channels. When senders implement
slow-start restart after a TCP connection goes idle (suggested by
Jacobson and Karels [JK92]), performance is reduced in long-lived
(but bursty) connections [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 or intermittent senders. 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).
3.7.2.2 Research
Simulation studies of rate-paced pacing for web-like traffic has
been shown to reduce 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].
3.7.2.3 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 to high for practical use.
Preliminary experiments suggest that in RBP this overhead is minimal
because RBP only requires this timer for the first RTT of
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transmission [VH98].
3.7.2.4 Topology Considerations
RBP could be used to restart an idle TCP connection 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.
3.8 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 1 byte 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.
3.8.1 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 [DNP97], [DENP97]
and [Jac90] all 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 these methods, the sender first sends a full TCP
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
information contained in later, compressed headers. This
compression can reduce the size of an IPv4/TCP header from 20 to as
few as 3 or 4 bytes.
Compression and decompression happen below the IP layer, and there
is a separate compressor / decompressor pair for each serial link.
Each compression pair maintains some state about some number of TCP
connections which may use the link concurrently, and the
decompresser passes 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.
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A variety of methods can be used by the endpoints of a connection to
negotiate the use of header compression. The PPP serial line
protocol allows for an option exchange, during which time the
endpoints can 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 it 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 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 it retransmits a packet with a full header in order to re-synch
the decompresser.
It is important to note that the compressor does not maintain any
timers, nor does the decompresser know when an error occured (only
the receiving TCP knows this, when the TCP checksum fails). A
single bit error will cause the decompresser to lose synch, 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 resynch 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.
[DNP97] and [DENP97] describe TCP/IPv4 and TCP/IPv6 compression
algorithms including compressing the various IPv6 extension headers
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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-synched. 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.
3.8.2 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.
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-synch the
decompresser in between 83% and 99% of the cases. This is
incredibly valuable, since packets received correctly after twice
has resynched the decompresser will cause duplicate acknowledgments.
This re-enables the use of both fast retransmit and SACK in
conjunction with header compression.
3.8.3 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.
3.8.4 Topology Considerations
TCP 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)
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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 become more critical.
3.9 Sharing TCP State Among Similar Connections
3.9.1 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 channels.
TCP includes a variety of parameters, many of which are set to
initial values which can severely affect the performance of
satellite connections, even though most TCP parameters are adjusted
later while the connection is established. These 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. It is difficult to select any single set of
parameters which is effective for all environments, however.
Instead of attempting to select these parameters a-priori, TCB
sharing keeps persistent state between incarnations of TCP
connections, and considers this state when initializing a new
connection. For example, if all connections to subnet 10 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.
Sharing state among connections brings up a number of questions such
as what 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. Finally, the security concerns associated
with sharing a piece of information need to be carefully considered
before introducing such a mechanism.
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3.9.2 Research
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; the point of TCB sharing is to raise this
to a research issue, and to further specify the ways in which the
information should be shared, regardless of the implementation.
3.9.3 Implementation Issues
Much of TCB sharing is an implementation issue only. The TCP
specifications do not preclude sharing information across
connections, or using some information from previous connections to
affect the state of new connections.
The goal of TCB sharing is to decouple the effect of connection
initialization from connection performance, to obviate the desire to
have persistent connections solely to maintain efficiency. This
allows separate connections to be more correctly used to indicate
separate associations, distinct from the performance implications
current implementations suffer.
Each TCP connection maintains state, usually in a data structure
called the TCP Control Block (TCB). The TCB contains information
about the connection state, its associated local process, and
feedback parameters about the connection's transmission. As
originally specified, and usually implemented, the TCB is maintained
on a per-connection basis. An alternate implementation can share
some of this state across similar connection instances and among
similar simultaneous connections. The resulting implementation can
have better transient performance, especially where long-term TCB
parameters differ widely from their typical initial values. These
changes can be constrained to affect only the TCB initialization,
and so have no effect on the long-term behavior of TCP after a
connection has been established. They can also be more broadly
applied to coordinate concurrent connections.
We note that the notion of sharing 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.
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Other implementation considerations are outlined in [Tou97] in
detail. Many instances of the implementation are the subject of
ongoing research.
3.9.4 Topology Considerations
TCB sharing aggregates state information. The set over which this
state is aggregated is critical to the performance of the
sharing. Worst case, nothing is shared, which degenerates to the
behavior of current implementations. Best case, information is
shared among connections sharing a critical property. In earlier
work [Tou97], the possibility of aggregating based on destination
subnet, or even routing path is considered.
For example, on a host connected to a satellite link, all
connections out of the host share the critical property of large
propagation latency, and are dominated by the bandwidth of the
satellite link. In this case, all connections with the same source
would share information.
It is expected that sharing state across TCP connections may be
useful in all network environments presented in section 2.
3.10 ACK Congestion Control
(in progress)
3.11 ACK Filtering
(in progress)
4 Mitigation Interactions
5 Conclusions
6 References
[AFP98] Sally Floyd, Mark Allman, Craig Partridge. Increasing TCP's
Initial Window, May 1997. Internet-Draft
draft-floyd-incr-init-win-03.txt (work in progress).
[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
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[AG98] Mark Allman, Dan Glover. Enhancing TCP Over Satellite
Channels using Standard Mechanisms, February 1998.
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[All97b] Mark Allman. Fixing Two BSD TCP Bugs. Technical Report
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[All98] Mark Allman. On the Generation and Use of TCP
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[AOK95] Mark Allman, Shawn Ostermann, Hans Kruse. Data Transfer
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[ASBD96] Vivek Arara, Narin Suphasindhu, John S. Baras, Douglas
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[BB95] Ajay Bakre, B.R. Badrinath. I-TCP: Indirect TCP for Mobile
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D. Estrin, S. Floyd, V. Jacobson, G. Minshall, C. Partridge,
L. Peterson, K. Ramakrishnan, S. Shenker, J. Wroclawski,
L. Zhang, Recommendations on Queue Management and Congestion
Avoidance in the Internet, April 1998. RFC 2309.
[BKVP97] B. Bakshi and P. Krishna and N. Vaidya and D. Pradham,
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H. Katz. The Effects of Asymmetry on TCP Performance. In
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[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.
[BPSK97] Hari Balakrishnan, Venkata N. Padmanabhan, Srinivasan
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TCP Performance over Wireless Links. IEEE/ACM Transactions on
Networking, December 1997.
[Bra89] Robert Braden. Requirements for Internet Hosts --
Communication Layers, October 1989. RFC 1122.
[Bra92] Robert Braden. Transaction TCP -- Concepts, September 1992.
RFC 1379.
[Bra94] Robert Braden. T/TCP -- TCP Extensions for Transactions:
Functional Specification, July 1994. RFC 1644.
[DENP97] Low-Loss TCP/IP Header Compression for Wirelesss Networks.
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", MOBICOMM 96, ACM, USA,
1996.
[DNP97] Mikael Degermark, Bjorn Nordgren, and Stephen Pink. IP
Header Compression, December 1997. Internet-Draft
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[FF98] Sally Floyd, Kevin Fall. Promoting the Use of End-to-End
Congestion Control in the Internet. Submitted to IEEE
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[FJ93] Sally Floyd and Van Jacobson. Random Early Detection
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Networking, V. 1 N. 4, August 1993.
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Computer Communication Review, V. 24 N. 5, October 1994.
[Hah94] Jonathan Hahn. MFTP: Recent Enhancements and Performance
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Center, June 1994.
[Hay97] Chris Hayes. Analyzing the Performance of New TCP
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[Hoe96] Janey Hoe. Improving the Startup Behavior of a Congestion
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[IL92] David Iannucci and John Lakashman. MFTP: Virtual TCP Window
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[Jac90] Van Jacobson. Compressing TCP/IP Headers, February 1990.
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RFC 1144.
[JBB92] Van Jacobson, Robert Braden, and David Borman. TCP
Extensions for High Performance, May 1992. RFC 1323.
[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
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[Kes91] Srinivasan Keshav. A Control Theoretic Approach to Flow
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[KAGT98] Hans Kruse, Mark Allman, Jim Griner, Diepchi Tran. HTTP
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[KM97] S. Keshav, S. Morgan. SMART Retransmission: Performance with
Overload and Random Losses. Proceeding of Infocom. 1997.
[MM96a] M. Mathis, J. Mahdavi, "Forward Acknowledgment: Refining TCP
Congestion Control," Proceedings of SIGCOMM'96, August, 1996,
Stanford, CA. Available from
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[MM96b] M. Mathis, J. Mahdavi, "TCP Rate-Halving with Bounding
Parameters" Available from
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[MSMO97] M. Mathis, J. Semke, J. Mahdavi, T. Ott, "The Macroscopic
Behavior of the TCP Congestion Avoidance Algorithm",Computer
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from Available from
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[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.
[PADHV97] Vern Paxson, Mark Allman, Scott Dawson, Ian Heavens,
Bernie Volz. Known TCP Implementation Problems, March 1998.
Internet-Draft draft-ietf-tcpimpl-prob-03.txt.
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[Par97] Craig Partridge. ACK Spacing for High Delay-Bandwidth Paths
with Insufficient Buffering, July 1997. Internet-Draft
draft-partridge-e2e-ackspacing-00.txt.
[Pax97] Vern Paxson. Automated Packet Trace Analysis of TCP
Implementations. In Proceedings of ACM SIGCOMM, September 1997.
[PN98] Poduri, K., and Nichols, K., Simulation Studies of Increased
Initial TCP Window Size, February 1998. Internet-Draft
draft-ietf-tcpimpl-poduri-00.txt (work in progress).
[Pos81] Jon Postel. Transmission Control Protocol, September 1981.
RFC 793.
[RF98] A Proposal to add Explicit Congestion Notification (ECN) to
IPv6 and to TCP. K. K. Ramakrishnan and Sally Floyd.
Internet-Draft draft-kksjf-ecn-01.txt, July 1998. (Work in
progress).
[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.
[SP97] Tim Shepard and Craig Partridge. When TCP Starts Up With
Four Packets Into Only Three Buffers, July 1997. Internet-Draft
draft-shepard-TCP-4-packets-3-buff-00.txt (work in progress).
[Ste97] W. Richard Stevens. TCP Slow Start, Congestion Avoidance,
Fast Retransmit, and Fast Recovery Algorithms, January 1997.
RFC 2001.
[Tou97] Touch, J., "TCP Control Block Interdependence," RFC-2140,
USC/Informatino Sciences Institute , 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.
7 Author's Addresses:
Mark Allman
NASA Lewis Research Center/Sterling Software
21000 Brookpark Rd. MS 54-2
Cleveland, OH 44135
mallman@lerc.nasa.gov
http://gigahertz.lerc.nasa.gov/~mallman
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Spencer Dawkins
Nortel
P.O.Box 833805
Richardson, TX 75083-3805
Spencer.Dawkins.sdawkins@nt.com
Dan Glover
NASA Lewis Research Center
21000 Brookpark Rd. MS 54-2
Cleveland, OH 44135
Daniel.R.Glover@lerc.nasa.gov
Jim Griner
NASA Lewis Research Center
21000 Brookpark Rd. MS 54-2
Cleveland, OH 44135
jgriner@lerc.nasa.gov
John Heidemann
University of Southern California/Information Sciences Institute
4676 Admiralty Way
Marina del Rey, CA 90292-6695
johnh@isi.edu
Shawn Ostermann
School of Electrical Engineering and Computer Science
Ohio University
416 Morton Hall
Athens, OH 45701
Phone: (740) 593-1234
ostermann@cs.ohiou.edu
Keith Scott
Jet Propulsion Laboratory
California Institute of Technology
4800 Oak Grove Drive MS 161-260
Pasadena, CA 91109-8099
Keith.Scott@jpl.nasa.gov
http://eis.jpl.nasa.gov/~kscott/
Jeffrey Semke
Pittsburgh Supercomputing Center
4400 Fifth Ave.
Pittsburgh, PA 15213
semke@psc.edu
http://www.psc.edu/~semke
Joe Touch
University of Southern California/Information Sciences Institute
4676 Admiralty Way
Marina del Rey, CA 90292-6695
USA
Phone: +1 310-822-1511 x151
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Fax: +1 310-823-6714
URL: http://www.isi.edu/~touch
Email: touch@isi.edu
Diepchi Tran
NASA Lewis Research Center
21000 Brookpark Rd. MS 54-2
Cleveland, OH 44135
dtran@lerc.nasa.gov
Expires: February, 1998 [Page 28]
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