Internet-Draft Prague Congestion Control October 2023
De Schepper, et al. Expires 16 April 2024 [Page]
Internet Congestion Control Research Group (ICCRG)
Intended Status:
K. De Schepper
Nokia Bell Labs
O. Tilmans
Nokia Bell Labs
B. Briscoe, Ed.
V. Goel
Apple Inc

Prague Congestion Control


This specification defines the Prague congestion control scheme, which is derived from DCTCP and adapted for Internet traffic by implementing the Prague L4S requirements. Over paths with L4S support at the bottleneck, it adapts the DCTCP mechanisms to achieve consistently low latency and full throughput. It is defined independently of any particular transport protocol or operating system, but notes are added that highlight issues specific to certain transports and OSs. It is mainly based on experience with the reference Linux implementation of TCP Prague and the Apple implementation over QUIC, but it includes experience from other implementations where available.

The implementation does not satisfy all the Prague requirements (yet) and the IETF might decide that certain requirements need to be relaxed as an outcome of the process of trying to satisfy them all. Future plans that have typically only been implemented as proof-of-concept code are outlined in a separate section.

Status of This Memo

This Internet-Draft is submitted in full conformance with the provisions of BCP 78 and BCP 79.

Internet-Drafts are working documents of the Internet Engineering Task Force (IETF). Note that other groups may also distribute working documents as Internet-Drafts. The list of current Internet-Drafts is at

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."

This Internet-Draft will expire on 16 April 2024.

1. Introduction

This document defines the Prague Congestion Control Algorithm (CCA), which is intended to maintain consistently low queuing delay over network paths that offer Low Latency, Low Loss, and Scalable throughput (L4S) support at the bottleneck [RFC9330]. Where the bottleneck does not support L4S, a Prague CCA is intended to fall back to behaving like a conventional 'Classic' congestion control. L4S support in the network involves Active Queue Management (AQM) that applies 'Immediate Explicit Congestion Notification (ECN)' at a very shallow target queueing delay (of the order of a millisecond) or even a virtual queue with no queuing delay, for example [RFC9332]. 'Immediate ECN' means that the network applies ECN marking based on the instantaneous queue, without any smoothing or filtering, The Prague CCA takes on the job of smoothing and filtering the congestion signals from the network.

The Prague CCA is a particular instance of a scalable congestion control, which is defined in Section 1.4. Scalable congestion control is the part of the L4S architecture that does the actual work of maintaining low queuing delay and ensuring that the delay and throughput properties scale with flow rate.

The L4S architecture [RFC9330] places the host congestion control in the context of the other parts of the system. In particular the different types of L4S AQM in the network and the codepoints in the IP-ECN field that convey to the network that the host supports the L4S form of ECN. The architecture document also covers other issues such as: incremental deployment; protection of low latency queues against accidental or malicious disruption; and the relationship of L4S to other low latency technologies. The specification of the L4S ECN Protocol [RFC9331] sets down the requirements that a Prague CCA has to follow (called the Prague L4S Requirements - see Section 2.1 for a summary).

This specification defines the Prague CCA independent of any particular transport protocol or operating system, but notes are added that highlight issues specific to certain transports and OSs. These primarily draw on experience with i) the reference implementation of Prague on Linux over TCP, and ii) the Apple implementation of Prague over QUIC. Nonetheless, wherever possible, experience from implementers on other platforms is included.

Links to implementations of the Prague CCA and other scalable congestion controls can be found via the L4S landing page [L4S-home], which also links to numerous other L4S-related resources. A (slightly dated) paper on the specific implementation of the Prague CCA in Linux over TCP is also available [PragueLinux], and the open-source code for Linux is at [linux-code].

1.1. Motivation: Low Queuing Delay /and/ Full Throughput

A Prague CCA is capable of keeping queuing delay consistently low while fully utilizing available capacity. In contrast, Classic congestion controls need to induce a reasonably large queue (approaching a bandwidth-delay product) in order to fully utilize capacity. Therefore, prior to scalable CCAs like DCTCP and Prague, it was believed that very low delay was only possible by limiting throughput and isolating the low delay traffic from capacity-seeking traffic.

A Prague CCA uses additive increase multiplicative decrease (AIMD), in which it increases its window until an ECN mark (or loss) is detected, then yields in a continual sawtooth pattern. The key to keeping queuing delay low without under-utilizing capacity is to keep the sawteeth tiny. This is achieved by ensuring that the reduction in rate that starts each sawtooth is proportionate to the prevailing level of congestion. In contrast, a classical CCA merely responds to the existence of congestion, not its extent. So each classical reduction has to be large enough to cope with the worst case.

For example the average duration of a Prague CCA sawtooth is of the order of a round trip, whereas a classic congestion control sawtooths over hundreds of round trips, implying multiple seconds. For instance, at 120 Mb/s and with a maximum RTT of 30 ms at the peak of each sawtooth, CUBIC takes 4.3 s to recover from each sawtooth reduction. At this rate, CUBIC is still fully in its Reno-friendly mode. If flow rate scales by 8 times to 960 Mb/s, it enters true CUBIC mode, with a recovery time of 12.2 s. From then on, each further scaling by 8 times doubles CUBIC's recovery time (because the cube root of 8 is 2), e.g., at 7.68 Gb/s, the recovery time is 24.3 s.

Keeping the sawtooth amplitude down keeps queue variation down and utilization up. Keeping the duration of the sawteeth down ensures control remains tight. The definition of a scalable CCA is that the duration between congestion marks does not increase as flow rate scales, all other factors being equal. This is important, because it means that the sawteeth will always stay tiny. So queue delay will remain very low, and control will remain very tight.

The tip of each sawtooth occurs when the bottleneck link emits a congestion signal. Therefore such small sawteeth are only feasible when ECN is used for the congestion signals. If loss were used after each brief recovery time, the loss level would be prohibitively high. This is why L4S-ECN has to depart from the requirement of Classic ECN[RFC3168] that an ECN mark is equivalent to a loss.

The Prague CCA is derived from Data Center TCP (DCTCP [RFC8257]). DCTCP is confined to controlled environments like data centres precisely because it uses such small sawteeth, which induce such a high level of congestion marking. For a CCA using Classic ECN, this would be interpreted as equivalent to the same, very high, loss level. The Classic CCA would then continually drive its own rate down in the face of such an apparently high level of congestion.

This is why coexistence with existing traffic is important for the Prague CCA. It has to be able to detect whether it is sharing a Classic bottleneck with Classic traffic, and if so fall back to behaving in a Classic way. If the bottleneck does not support ECN at all, that is easy - a Prague CCA just responds in the Classic way to loss (see Section 2.4.1). But if it is sharing a Classic ECN bottleneck with Classic ECN traffic, this is more difficult to detect (see Section 3.3). Because the Prague CCA removes most of the queue, it also has to address RTT-dependence. Otherwise, at low base RTTs, its flow rate would become far more RTT-dependent than Classic CCAs.

1.2. Document Purpose

There is not 'One True Prague CCA'. L4S is intended to enable development of any scalable CCA that meets the L4S Prague requirements [RFC9331]. This document attempts to describe a design that transfers a byte stream. It is generalized across different transports and OS platforms.

1.3. Maturity Status (To be Removed Before Publication)

The field of congestion control is always a work in progress. However, there are areas of Prague implementations that are still just placeholders while separate research code is evaluated. And in other implementations of the Prague CCA, other areas are incomplete. In the Linux reference implementation of TCP Prague, interim code is used in the incomplete areas, which are:

  • Flow start and restart (standard slow start is used, even though it often exits early in L4S environments were ECN marking tends to be frequent);

  • Faster than additive increase (standard Reno additive increase is used, which makes the flow particularly sluggish if it has dropped out of slow start early).

The body of this document describes the Prague CCA as implemented. Any non-default options or any planned improvements are separated out into Section 3 on "Variants and Future Work". As each of the above areas is addressed, it will will be removed from this section and its description in the body of the document will be updated. Once all areas are complete, this section will be removed. Prague will then still be a work in progress, but only on a similar footing as all other congestion controls.

No implementation satisfies all the Prague requirements (yet), and the IETF might decide that certain requirements need to be relaxed as an outcome of the process of trying to satisfy them all.

1.4. Terminology

The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT", "SHOULD", "SHOULD NOT", "RECOMMENDED", "NOT RECOMMENDED", "MAY", and "OPTIONAL" in this document are to be interpreted as described in BCP 14 [RFC2119] [RFC8174] when, and only when, they appear in all capitals, as shown here.

Definitions of terms:

Classic Congestion Control:
A congestion control behaviour that can co-exist with standard TCP Reno [RFC5681] without causing significantly negative impact on its flow rate [RFC5033]. With Classic congestion controls, as flow rate scales, the number of round trips between congestion signals (losses or ECN marks) rises with the flow rate. So it takes longer and longer to recover after each congestion event. Therefore control of queuing and utilization becomes very slack, and the slightest disturbance prevents a high rate from being attained [RFC3649].
Scalable Congestion Control:
A CCA where the average time from one congestion signal to the next (the recovery time) remains invariant as the flow rate scales, all other factors being equal. This maintains the same degree of control over queueing and utilization whatever the flow rate, as well as ensuring that high throughput is robust to disturbances. For instance, DCTCP averages 2 congestion signals per round-trip whatever the flow rate. For the public Internet a Scalable CCA has to comply with the requirements in Section 4 of [RFC9331] (aka. the 'Prague L4S requirements').
Response function:
The relationship in steady state between the window (cwnd) of a congestion control and the congestion signalling probability, p, [RFC3649]. A widely used family of response functions has the form cwnd = K/p^B, where K and B are constants. In an approximation of the response function of the standard Reno CCA, B=1/2. For a scalable congestion control B=1, so its response function takes the form cwnd = K/p. The number of congestion signals per round is p*cwnd, which equates to the constant, K, for a scalable CCA. Hence the definition of a scalable CCA above.
The subset of Classic traffic that excludes unresponsive traffic and excludes experimental congestion controls intended to coexist with Reno but without always being strictly friendly to it (as allowed by [RFC5033]). Reno-friendly is used in place of 'TCP-friendly', given that the TCP protocol is used with many different congestion control behaviours.
Classic ECN:

The original Explicit Congestion Notification (ECN) protocol [RFC3168], which requires ECN signals to be treated the same as drops, both when generated in the network and when responded to by the sender.

The names used for the four codepoints of the 2-bit IP-ECN field are as defined in [RFC3168]: Not ECT, ECT(0), ECT(1) and CE, where ECT stands for ECN-Capable Transport and CE stands for Congestion Experienced.

A packet marked with the CE codepoint is termed 'ECN-marked' or sometimes just 'marked' where the context makes ECN obvious.

Congestion Control
Congestion Control Algorithm
an ACKnowledgement, or to ACKnowledge
Exponentially Weighted Moving Average
Round Trip Time

Definitions of Parameters and Variables:

Maximum transmission unit [b]
Congestion window [B]
Slow start threshold [B]
The amount of data that the sender has sent but not yet received ACKs for [B]
Steady-state probability of drop or marking []
EWMA of the ECN marking fraction []
the amount of new data acknowledged by an ACK [B]
the amount of newly acknowledged data that was ECN-marked [B]
additive increase to apply per RTT [B]
Smoothed round trip time [s]
Maximum allowed bottleneck queuing delay due to segmentation offload bursts [s] (default 0.25 ms for the public Internet)

2. Prague Congestion Control

2.1. The Prague L4S Requirements

The beneficial properties of L4S traffic (low queuing delay, etc.) depend on all L4S sources satisfying a set of conditions called the Prague L4S Requirements. The name is after an ad hoc meeting of about thirty people co-located with the IETF in Prague in July 2015, the day after the first public demonstration of L4S.

The meeting agreed on a list of modifications to DCTCP [RFC8257] to focus activity on a variant that would be safe to use over the public Internet. It was suggested that this could be called TCP Prague to distinguish it from DCTCP. This list was adopted by the IETF, and has continued to evolve (see section 4 of [RFC9331]). The requirements are no longer TCP-specific, applying irrespective of wire-protocol (TCP, QUIC, RTP, SCTP, etc).

This unusual start to the life of the project led to the unusual development process of a reference implementation that had to resolve a number of ambitious requirements, already known to be in tension [Tensions17].

DCTCP already implements a scalable congestion control. So most of the changes to make it usable over the Internet seemed trivial, some 'merely' involving adoption of other parallel developments like Accurate ECN TCP feedback [I-D.ietf-tcpm-accurate-ecn] or RACK [RFC8985]. Others have been more challenging (e.g. RTT-independence). And others that seemed trivial became challenging given the complex set of bugs and behaviours that characterize today's Internet and a modern end system stack such as Linux.

The more critical implementation challenges are highlighted in the following sections, in the hope we can prevent mistakes being repeated (see for instance Section 2.3.2, Section 2.4.2). There was also a set of five intertwined 'bugs' - all masking each other, but causing unpredictable or poor performance as different code modifications or environments unmasked them. A comprehensive write-up of the experiments that had to be designed to isolate these bugs is available separately [DCTCP_flaws].

During the development process, we have unearthed fundamental aspects of the implementation and indeed the design of DCTCP and Prague that have still not caught up with the paradigm shift from existence to extent-based congestion response. Some have been implemented by default, e.g. not suppressing additive increase for a round trip after a congestion event (Section 2.4.3). Others have been implemented but not fully evaluated, e.g. removing the 1-2 unnecessary round trips of lag in feedback processing (Section 3.1.3) and yet others are still future plans, e.g. further RTT-independence (Section 3.4) and exploiting combined congestion metrics in more cases (Section 3.2).

The requirements are categorized into those that would impact other flows if not handled properly and performance optimizations that are important but optional from the IETF's point of view, because they only affect the flow itself. The list below maps the order of the requirements in [RFC9331] to the order in this document (which is by functional categories and code status):

Mandatory or Advisory Requirements:
Optional performance optimizations:

2.2. Packet Identification

On the public Internet, a sender using a Prague CCA MUST set the ECT(1) codepoint on all the packets it sends, in order to identify itself as an L4S-capable congestion control (Req 4.1. [RFC9331]).

This applies whatever the transport protocol, whether TCP, QUIC, RTP, etc. In the case of TCP, unlike an RFC 3168 TCP ECN transport, a sender can set all packets as ECN-capable, including TCP control packets and retransmissions [RFC8311], [I-D.ietf-tcpm-generalized-ecn].

A Prague CCA SHOULD optionally be configurable to use the ECT(0) codepoint in private networks, such as data centres, which might be necessary for backward compatibility with DCTCP deployments where ECT(1) might already have another usage.

Linux implementation note:

TCP Prague in Linux kernel:
The Linux kernel was updated to allow the ECT(1) flag to be set from within a CCA module. A Prague CCA then has full control over the ECN code point it uses at any one time. In this way it asserts the use of ECT(1) (or optionally ECT(0)) and non-ECT when required.

2.3. Detecting and Measuring Congestion

2.3.1. Accurate ECN Feedback

When feedback of ECN markings was added to TCP [RFC3168], it was decided not to report any more than one mark per RTT. L4S-capable congestion controls need to know the extent, not just the existence of congestion (Req 4.2. [RFC9331]). Recently defined transports (DCCP, QUIC, etc) typically already satisfy this requirement. So they are dealt with separately below, while TCP and derivatives such as SCTP [RFC9260] are covered first. Accurate ECN Feedback with TCP & Derivatives

The TCP wire protocol is being updated to allow more accurate feedback (AccECN [I-D.ietf-tcpm-accurate-ecn]). Therefore, in the case where a sender uses a Prague CCA over TCP, whether as client or server:

  • it MUST itself support AccECN;

  • to support AccECN it also has to check that its peer supports AccECN during the handshake.

If the peer does not support accurate ECN feedback, the sender MUST fall back to a Reno-friendly CCA behaviour for the rest of the connection. The non-Prague TCP sender MUST then no longer set ECT(1) on the packets it sends. Note that the peer only needs to support AccECN; there is no need to find out whether the peer is using an L4S-capable congestion control.

Note that a sending TCP client that uses a Prague CCA can set ECT(1) on the SYN prior to checking whether the other peer supports AccECN (as long as it follows the procedure in [I-D.ietf-tcpm-generalized-ecn] if it subsequently discovers the peer does not support AccECN).

Linux implementation note:

TCP Prague in Linux kernel:

The Linux kernel update to support AccECN for TCP is independent of the CCA module in use. So the kernel tries to negotiate AccECN feedback whichever congestion control module is selected. The Prague CCA module does an additional check to verify that the kernel actually does support AccECN, based on which Prague will either decide to proceed using a scalable CCA or fall back to a Classic CCA (Reno in the current Linux implementation).

A system-wide option is available to enable or disable AccECN negotiation. However, TCP flows using the Prague CCA module depend on AccECN; so they always ignore this system-wide sysctl and enable AccECN negotiation anyway. Accurate ECN Feedback with Other Modern Transports

Transport protocols specified recently, .e.g. DCCP [RFC4340], QUIC [RFC9000], are unambiguously suitable for Prague CCAs, at least according to the protocol specs, because they were designed from the start with accurate ECN feedback. However, in practice, ECN feedback was omitted from the early versions of some QUIC implementations.

At the start of a connection, a QUIC sender that wishes to use ECN checks incoming feedback for whether ECN correctly traversed the forward path. To such a QUIC sender, a peer that has not implemented ECN feedback is indistinguishable from a network path that bleaches the ECN field. So, in either case, the sender will disable the use of ECN.

Without ECN, Prague essentially turns into a Classic CCA. For instance, the Apple implementation of Prague for QUIC responds to each loss as CUBIC would. So without ECN, its behaviour becomes indistinguishable from CUBIC. This is because Prague's response to each loss is Reno-friendly (as required in Section 2.4.1).

In the case of RTP/RTCP, ECN feedback was added in [RFC6679], which is sufficient for a Prague CCA. However, it is preferable to use the most recent improvements to ECN feedback in [RFC8888], as used in the implementation of the L4S variant of SCReAM [RFC8298].

2.3.2. Moving Average of ECN Feedback

A Prague CCA currently maintains a moving average of ECN feedback in a similar way to DCTCP. This section is provided mainly because performance has proved to be sensitive to implementation precision in this area. So first, some background is necessary.

A Prague CCA triggers update of its moving average once per rtt_virt (see Section 2.4.4). To maintain its moving average, it measures the fraction, frac, of ACKed packets {ToDo: add Future Work on using ACKed bytes instead, if decided appropriate} that carried ECN feedback over the previous rtt_virt. It then updates an exponentially weighted moving average (EWMA) of this fraction, called alpha, using the following algorithm:

  • alpha += g * (frac - alpha);

where g (default 1/16) is the gain of the EWMA.

The moving average, alpha, is initialized to 1 at the first sign of ECN feedback, which ensures the maximum congestion response to the first appearance of congestion at a bottleneck supporting ECN.

Linux implementation notes:

Rounding problems in DCTCP:
Alpha is a fraction between 0 and 1, and it needs to be represented with high resolution because the larger the bandwidth-delay product (BDP) of a flow, the smaller the value that alpha converges to (in steady state alpha = 2/cwnd). In principle, Linux DCTCP maintains the moving average 'alpha' using the same formula as Prague CCA uses (as above). Linux DCTCP represents alpha with a 10-bit integer (with resolution 1/1024). However, up to kernel release 3.19, Linux used integer arithmetic that could not reduce alpha below 15/1024. Then it was patched so that any value below 16/1024 was rounded down to zero [patch-alpha-zero]. For a flow with a higher BDP than 128 segments, this means that, alpha flip-flops. Once it has flopped down to zero DCTCP becomes unresponsive until it has built sufficient queue to flip up to 16/1024. For larger BDPs, this causes DCTCP to induce larger sawteeth, which loses the low-queuing-delay and high-utilization intent of the algorithm.
Upscaled alpha in Prague CC:

To resolve the above problem the implementation of TCP Prague in Linux maintains upscaled_alpha = alpha/g instead of alpha:

  • upscaled_alpha += frac - g * upscaled_alpha;

This technique is the same as Linux uses for the retransmission timer variables, srtt and mdev. The Linux Prague CCA also uses 20 bits for alpha.

The above per-rtt_virt update to the moving average is currently used in the Linux Prague CCA. However, another approach is being investigated because these per-RTT updates introduce 1--2 rounds of delay into the congestion response on top of the inherent round of feedback delay (see Section 3.1.3 in the section on variants and future work).

2.3.3. Scaling Loss Detection with Flow Rate

After an ACK leaves a gap in the sequence space, a Prague CCA is meant to deem that a loss has occurred using 'time-based units' (Req 4.3. [RFC9331]). This is in contrast to the traditional approach that counts a hard-coded number of duplicate ACKs, e.g. the 3 Dup-ACKs specified in [RFC5681]. The reason is that counting packets rather than time unnecessarily tightens the time within which parallelized links have to keep packets in sequence as flow rate scales over the years.

To satsify this requirement, a Prague CCA SHOULD wait until a certain fraction of the RTT has elapsed before it deems that the gap is due to packet loss. The reference implementation of TCP Prague in Linux uses RACK [RFC8985] to address this requirement. An approach similar to TCP RACK is also used in QUIC.

At the start of a connection, RACK counts 3 DupACKs to detect loss because the initial smoothed RTT estimate can be inaccurate. This would depend indirectly on time as long as the initial window (IW) is paced over a round trip (see Section 2.4.5). For instance, if the initial window of 10 segments was paced evenly across the initial RTT then, in the next round, an implementation that deems there has been a loss after (say) 1/4 of an RTT can count 1/4 of 10 = 3 DupACKs (rounded up). In subsequent rounds, if the TCP connection detects reordering, RACK shifts to using a fraction of the RTT for loss detection.

2.4. Congestion Response Algorithm

In congestion avoidance phase, a Prague CCA uses a similar additive increase multiplicative decrease (AIMD) algorithm to DCTCP, but with the following differences:

2.4.1. Loss behaviour

A Prague CCA MUST use a Reno-friendly congestion response (such as that of CUBIC [RFC9438] or Reno [RFC5681]) on detection of a loss (Req 2 in section 4.3. of [RFC9331]). DCTCP falls back to Reno for the round trip after a loss, and the Linux reference implementation of TCP Prague currently inherits this behaviour. On detection of loss, the Apple implementation of Prague uses CUBIC's behaviour instead of Reno's, for both the reduction after the loss and the subsequent growth of cwnd until the next congestion event.

If a Prague CCA has already reduced the congestion window due to ECN feedback less than a round trip before it detects a loss, it MAY reduce the congestion window by a smaller amount due to the loss, as long as the reductions due to ECN and the loss, when combined, result in a reduction that is at least as large as the implementation usually makes in response to loss (e.g. 50% to emulate Reno or 30% to emulate CUBIC).

See Section 3.2 for discussion of future work on congestion control using a combination of delay, ECN and loss.

Linux implementation note:

DCTCP bug prior to v5.1:
A Prague CCA cannot rely on inheriting the fall-back-on-loss behaviour of the DCTCP code in the Linux kernel prior to v5.1, due to a previous bug in the fast retransmit code (but not in the retransmission timeout code) [patch-loss-react].

2.4.2. Multiplicative Decrease on ECN Feedback

A Prague CCA currently responds to ECN feedback in a similar way to DCTCP. This section is provided mainly because performance has proved to be sensitive to implementation details in this area. So the following recap of the congestion response is needed first.

As explained in Section 2.3.2, like DCTCP, the Linux Prague CCA clocks its moving average of ECN-marking, alpha, once per rtt_virt throughout a connection. Nonetheless, it only triggers a multiplicative decrease to its congestion window when it actually receives an ACK carrying ECN feedback. Then it suppresses any further decreases for one round trip, even if it receives further ECN feedback. This is termed Congestion Window Reduced or CWR state.

A Prague CCA (also like DCTCP) ensures that the average recovery time remains invariant as flow rate scales (Req 4.3 of [RFC9331]) by making the multiplicative decrease depend on the prevailing value of alpha as follows:

  • ssthresh = (1 - alpha/2) * cwnd;

Linux implementation notes:

Upscaled alpha:
With reference to the earlier discussion of integer arithmetic precision (Section 2.3.2), alpha = g * upscaled_alpha.
Carry of fractional cwnd remainder:

Typically the absolute reduction in the window is only a small number of segments. So, if the Prague CCA implementation counts the window in integer segments (as in the Linux reference code), delay can be made significantly less jumpy by tracking a fractional value, cwnd_carry, alongside the integer window and carrying over any fractional remainder to the next reduction. Also, integer rounding bias ought to be removed from the multiplicative decrease calculation. Specifically:

#define ONE_CWND (1LL << 20)        /* Must be signed */
#define MAX_ALPHA (1ULL << 20)

/* On CE feedback, calculate the reduction in cwnd */
    /* Adding MAX_ALPHA to the numerator effectively adds 1/2
     *  which compensates for integer division rounding down */
    reduction = (alpha*cwnd*ONE_CWND + MAX_ALPHA) / MAX_ALPHA / 2;
    cwnd_carry -= reduction;

/* Round reduction into whole segments and carry the remainder */
    if (cwnd_carry <= -ONE_CWND) {
        cwnd_carry += ONE_CWND;
        cwnd = max(cwnd - 1, MIN_CWND);
        ssthresh = cwnd;

In dynamic scenarios, as flows find a new operating point, alpha will have often tailed away to near-nothing before the onset of congestion. Then DCTCP's tiny reduction followed by no further response for a round is precisely the wrong way for a CCA to respond. A solution to this problem is being evaluated as part of the work already mentioned to improve Prague's responsiveness (see Section 3.1.3 in the section on variants and future work).

2.4.3. Additive Increase and ECN Feedback

Unlike DCTCP, a Prague CCA does not suppress additive increase for one round trip after a congestion window reduction (i.e., while in CWR state). Instead, a Prague CCA applies additive increase irrespective of its CWR state, but only for bytes that have been ACK'd without ECN feedback. Specifically, on each ACK,

    cwnd += (acked_sacked - ece_delta) * ai_per_rtt / cwnd;


  • acked_sacked is the number of new bytes acknowledged by the ACK;

  • ece_delta is the number of newly acknowledge ECN-marked bytes;

  • ai_per_rtt is a scaling factor that is typically 1 SMSS except for small RTTs (see Section 2.4.4)

Superficially, the classical suppression of additive increase for the round after a decrease seems to make sense. However, DCTCP and Prague are designed to induce an average of 2 congestion marks per RTT in steady state, which leaves very little space for any increase between the end of one round of CWR and the next mark. In tests, when a test version of a Prague CCA is configured to completely suppress additive increase during CWR (like Reno and DCTCP), its sawteeth become more irregular, which is its way of making some decreases large enough to open up enough space for an increase. This irregularity tends to reduce link utilization. Therefore, the Prague CCA has been modified to continue additive increase irrespective of CWR state.

Nonetheless, rather than continue additive increase regardless of congestion, it is safer to only increase on those ACKs that do not feed back congestion. This approach reduces additive increase as the marking probability increases, which tends to keep the marking level unsaturated (below 100%) (see Section 3.1 of [Tensions17]). Under stable conditions, Prague's congestion window then becomes proportional to (1-p)/p, rather than 1/p.

See also 'Faster than Additive Increase' (Section 3.1.2)

2.4.4. Reduced RTT-Dependence

The window-based AIMD described so far was inherited from Reno via DCTCP. When many long-running Reno flows share a link, their relative packet rates become roughly inversely proportional to RTT. Then a flow with very small RTT will dominate any flows with larger RTTs.

Queuing delay sets a lower limit to the smallest possible RTT. So, prior to the extremely low queuing delay of L4S, extreme cases of RTT dependence had never been apparent. Now that L4S has removed most of the queuing delay, we have to address the root-cause of RTT-dependence, which a Prague CCA is required to do, at least when the RTT is small (see the 'Reduced RTT bias' aspect of Req 4.3. [RFC9331]). Here, a small RTT is defined as below "typical RTTs expected in the intended deployment scenario".

The Linux reference Prague CCA reduces RTT bias by using a virtual RTT (rtt_virt) rather than the actual smoothed RTT (srtt) for all three of: i) the period of additive window increase; ii) the EWMA update period; and iii) the duration of CWR state after a decrease. rtt_virt is calculated as a function of the actual smoothed RTT, chosen so that, when the srtt is high, the virtual RTT is essentially the same; but for lower actual RTTs, the virtual RTT is increasingly larger than the actual RTT. Example functions that implementations might use for the virtual RTT are:

  • rtt_virt = max(srtt, RTT_VIRT_MIN);

  • rtt_virt = srtt + AdditionalRTT;

where the parameters RTT_VIRT_MIN or AdditionalRTT would be set for a particular deployment environment.

The Linux implementation of Prague uses the first example and for the public Internet sets RTT_VIRT_MIN=25ms. Thus, Linux Prague defines rtt_virt = max(srtt, 25ms), which addresses the Prague requirement for Reduced RTT-Dependence when the RTT is smaller than typical public Internet RTTs.

As the actual window (cwnd) is still sent within 1 actual RTT, we also need to use a (conceptual) virtual window, cwnd_virt. For instance, if rtt_virt = 25 ms then, when the actual RTT is 5 ms, there are rtt_virt/srtt = 5 times more packets in cwnd_virt, than in the actual window, cwnd, because cwnd_virt spans 5 actual round trips. We define M as the ratio rtt_virt/srtt.

In Reno or DCTCP, additive increase is implemented by dividing the desired increase of 1 segment per round over the cwnd packets in the round. This requires an increase of 1/cwnd per packet.

In the Linux implementation of TCP Prague, the aim is to increase the reference window by 1 segment over a virtual RTT. However, in practice the increase is applied to the actual window, cwnd, which is M times smaller than cwnd_virt. So cwnd has to be increased by only 1/M segments over rtt_virt. But again, in practice, the increase is applied over an actual window of packets spanning an actual RTT, which is also M times smaller than the virtual RTT. So the desired increase in cwnd is only 1/M^2 segments over an actual round trip containing cwnd packets. Therefore, the increase in cwnd per packet has to be (1/M^2) * (1/cwnd).

Unless a flow lasts long enough for rates to converge, aiming for equal rates will not be relevant. So, in the Linux implementation of Prague, the Reduced RTT-Dependence algorithm only comes into effect after D rounds, where D is configurable (current default 500). Continuing the previous example, if actual srtt=5 ms and rtt_virt = 25 ms, Prague would use the regular RTT-dependent algorithm for the first 500*5ms = 2.5s. Then it would start to converge to more equal rates using its Reduced RTT-Dependence algorithm. If the actual RTT were higher (e.g. 20ms), it would stay in the regular RTT-dependent mode for longer (500 rounds = 10s), but this would be mitigated by the actual RTT it uses at the start being closer to the virtual RTT that it eventually uses (20ms and 25ms resp.).

This approach prevents reduced RTT-dependence from making the flow less responsive at start-up and ensures that its early throughput share is based on its actual RTT. The benefit is that short flows (mice) give themselves priority over longer flows (elephants), and shorter RTTs will still converge faster than longer RTTs. Nonetheless, the throughput still converges to equal rates after D rounds.

It is planned to reset the algorithm to the regular RTT-dependent behaviour after an idle, not just at flow start, as discussed under Future Work in Section 3.4.

Section 3.4 also discusses extending the reduction in RTT-dependence to longer RTTs than RTT_VIRT_MIN (i.e. longer than 25ms). The current Linux Prague implementation does not support this.

2.4.5. Flow Start or Restart

Currently the Linux reference implementation of TCP Prague uses the standard Linux slow start code. Slow start is exited once a single mark is detected.

When other flows are actively filling the link, regular marks are expected, causing slow start of new flows to end prematurely. This is clearly not ideal, so other approaches are being worked on (see Section 3.1.1). However, slow start has been left as the default until a properly matured solution is completed.

2.5. Packet Sending

2.5.1. Packet Pacing

A Prague CCA MUST pace the packets it sends to avoid the queuing delay and under-utilization that would otherwise be caused by bursts of packets that can occur, for example, when restarting from idle or a jump in the acknowledgement number opens up cwnd. Prague does this in a similar way to the base Linux TCP stack, by spacing out the window of packets evenly over the round trip time, using the following calculation of the pacing rate [b/s]:

  • pacing_rate = MTU_BITS * max(cwnd, inflight) / srtt;

During slow start, as in the base Linux TCP stack, Prague factors up pacing_rate by 2, so that it paces out packets twice as fast as they are acknowledged. This keeps up with the doubling of cwnd, but still prevents bursts in response to any larger transient jumps in cwnd.

    if (cwnd < ssthresh / 2)
        pacing_rate *= 2;

During congestion avoidance, the Linux TCP Prague implementation does not factor up pacing_rate at all. This contrasts with the base Linux TCP stack, which currently factors up pacing_rate by a ratio parameter set to 1.2. The developers of the base Linux stack confirmed that this factor of 1.2 was only introduced in case it improved performance, but there were no scenarios where it was known to be needed. In testing of Prague, this factor was found to cause queue delay spikes whenever cwnd jumped more than usual. And throughput was no worse without it. So it was removed from the Linux TCP Prague CCA.

A Prague CCA can use alternatives to the traditional slow-start algorithm that use different pacing (see Section 2.4.5).

2.5.2. Segmentation Offload

In the absence of hardware pacing, it becomes increasingly difficult for a machine to scale to higher flow rates unless it is allowed to send packets in larger bursts, for instance using segmentation offload. Happily, as flow rate scales up, proportionately more packets can be allowed in a burst for the same amount of queuing delay at the bottleneck.

Therefore, a Prague CCA sends packets in a burst as long as it will not induce more than MAX_BURST_DELAY of queuing at the bottleneck. From this constant and the current pacing_rate, it calculates how many MTU-sized packets to allow in a burst:

  • max_burst = pacing_rate * MAX_BURST_DELAY / MTU_BITS

The current default in the Linux TCP Prague for MAX_BURST_DELAY is 250us which supports marking thresholds starting from about 500us without underutilization. This approach is similar to that in the Linux TCP stack, except there MAX_BURST_DELAY is 1ms.

3. Variants and Future Work

3.1. Getting up to Speed Faster

Appendix A.2. of [RFC9331] outlines the performance optimizations needed when transplanting DCTCP from a DC environment to a wide area network. The following subsections address two of those points: faster flow startup and faster than additive increase. Then Section 3.1.3 covers the flip side, in which established flows have to yield faster to make room, otherwise queuing will result.

3.1.1. Flow Start (or Restart)

For faster flow start, two approaches are currently being investigated in parallel:

Modified Slow Start:

The traditional exponential slow start can be modified both at the start and the end, with the aim of reducing the risk of queuing due to bursts and overshoot:

Pacing IW:

A Prague CCA can use an initial window of 10 (IW10 [RFC6928]), but pacing of this Initial Window is recommended to try to avoid the pulse of queuing that could otherwise occur. Pacing IW10 also spreads the ACKs over the round trip so that subsequent rounds consist of ten subsets of packets (with 2, 4, 8 etc. per round in each subset), rather than a single set with 20, 40, 80 etc. in each round. With IW paced, if a queue builds during a round (e.g. due to other unexpected traffic arriving) it can drain in the gap before the next subset, rather than the whole set backing up into a much larger queue. As smoothed RTT is unknown or inaccurate at the start of a flow, an implementation can pace IW over a fraction of the initial smoothed RTT (perhaps also clamped between hard-coded sanity limits). The implementation could also initialize SRTT with a value it had previously cached per destination (as long as it is sufficiently fresh). The safety factor could depend on whether a cached value was available and how recent it was.

In the Linux reference implementation of TCP Prague, IW pacing can be optionally enabled, but it is off by default, because it is yet to be fully evaluated. It currently paces IW over half the initial smoothed round trip time (SRTT) measured during the handshake. SRTT is halved because the RTT often reduces after the initial handshake. For example: i) some CDNs move the flow to a closer server after establishment; ii) the initial RTT from a server can include the time to wake a sleeping handset radio; iii) some uplink technologies take a link-level round trip to request a scheduling slot.

It is also planned to exploit any cached knowledge of the path RTT to improve the initial estimate, for instance using the Linux per-destination cache. It is also planned to allow the application to give an RTT hint (by setting sk_max_pacing_rate in Linux) if the developer has reason to believe that the application has a better estimate.

Exiting slow start more gracefully:
In the wide area Internet (in contrast to data centres), bottleneck access links tend to have much less capacity than the line rate of the sender. With a shallow immediate ECN threshold at this bottleneck, the slightest burst can tend to induce an ECN mark, which traditionally causes slow start to exit. A more gradual exit is being investigated for a Prague CCA using the extent of marking, not just the existence of a single mark. This will be more consistent with the extent-based marking that scalable congestion controls use during congestion avoidance. Delay measurements (similar to Hystart++ [RFC9406]) can also be used to complement the ECN signals.
Paced Chirping:

In this approach, the aim is to both increase more rapidly than exponential slow-start and to greatly reduce any overshoot. It is primarily a delay-based approach, but the aim is also to exploit ECN signals when present (while not forgetting loss either). Therefore Paced Chirping is generally usable for any congestion control - not solely for a Prague CCA and L4S.

Instead of only aiming to detect capacity overshoot at the end of flow-start, brief trains of rapidly decreasing inter-packet spacing called chirps are used to test many rates with as few packets and as little load as possible. A full description is beyond the scope of this document. [LinuxPacedChirping] introduces the concepts and the code as well as citing the main papers on Paced Chirping.

Paced chirping works well over continuous links such as Ethernet and DSL. But better averaging and noise filtering are necessary over discontinuous link technologies such as WiFi, LTE cellular radio, passive optical networks (PON) and data over cable (DOCSIS). This is the current focus of this work.

The current Linux implementation of TCP Prague does not include Paced Chirping, but research code is available separately in Linux and ns3. it is accessible via the L4S landing page [L4S-home].

3.1.2. Faster than Additive Increase

A Prague CCA has a startup phase and congestion avoidance phase like traditional CCAs. In steady-state during congestion avoidance, like all scalable congestion controls, it induces frequent ECN marks, with the same average recovery time between ECN marks, no matter how much the flow rate scales.

If available capacity suddenly increases, e.g. other flow(s) depart or the link capacity increases, these regular ECN marks will stop. Therefore after a few rounds of silence (no ECN marks) in congestion avoidance phase, a Prague CCA can assume that available capacity has increased, and switch to using the techniques from its startup phase (Section 3.1.1) to rapidly find the new, faster operating point. Then it can shift back into its congestion avoidance behaviour.

That is the theory. But, as explained in Section 3.1.1, the startup techniques, specifically paced chirping, are still being developed for discontinuous link types. Once the startup behaviour is available, the Linux implementation of a Prague CCA will also have a faster than additive increase behaviour. S.3.2.3 of [PragueLinux]) gives a brief preview of the performance of this approach over an Ethernet link type in ns3.

3.1.3. Remove Lag in Congestion Response

To keep queuing delay low, new flows can only push in fast if established flows yield fast. It has recently been realized that the design of a Prague EWMA and congestion response introduces 1-2 rounds of lag (on top of the inherent round of feedback delay due to the speed of light). These lags were inherited from the design of DCTCP (see Section 2.3.2 and Section 2.4.2), where a couple of extra hundred microseconds was less noticeable. But congestion control in the wide area Internet cannot afford up to 2 rounds trips of extra lag.

To be clear, lag means delay before any response at all starts. That is qualititatively different from the smoothing gain of an EWMA, which /reduces/ the response by the gain factor (1/16 by default) in case a change in congestion does not persist. Smoothing gain can always be increased. But 1-2 rounds of lag means that, when a new flow tries to push in, the sender of an established flow will not respond /at all/ for 1-2 rounds after it first receives congestion feedback.

A Prague CCA spends the first round trip of this lag gathering feedback to measure frac before it is input into the EWMA algorithm (see Section 2.3.2). Then there is up to one further round of delay because the implementations of DCTCP and Prague did not fully adopt the paradigm shift to extent-based marking - the timing of the decrease is still based on Reno.

Both Reno and DCTCP/Prague respond immediately on the first sign of congestion. Reno's response is large, so it waits a round in CWR state to allow the response to take effect. DCTCP's response is tiny (extent-based), but then it still waits a round in CWR state. So it does next-to-nothing for a round.

New EWMA and response algorithms to remove these 1-2 extra rounds of lag are described in [PerAckEWMA]. They have been implemented in Linux and an iterative process of evaluation and redesign is in progress. The EWMA is updated per-ACK, but it still changes as if it is clocked per round trip. The congestion response is still triggered by the first indication of ECN feedback, but it proceeds over the subsequent round trip so that it can take into account further incoming feedback as the EWMA evolves. The reduction is applied per-ACK but sized to result as if it had been a single response per round trip.

3.2. Combining Congestion Metrics

Ultimately, it would be preferable to take an integrated approach and use a combination of ECN, loss and delay metrics to drive congestion control. For instance, using a downward trend in ECN marking and/or delay as a heuristic to temper the response to loss. Such ideas are not in the immediate plans for the Linux TCP Prague, but some more specific ideas are highlighted in the following subsections.

3.2.1. ECN with Loss

If the bottleneck is ECN-capable, a loss due to congestion is very likely to have been preceded by a period of ECN marking. When the current Linux TCP Prague CCA detects a loss, like DCTCP, it halves cwnd, even if it has already reduced cwnd in the same round trip due to ECN marking. This double reduction can end up factoring down cwnd to as little as 1/4 in one round trip. This is not necessarily detrimental (experimentation will tell), but, if necessary, the response to loss can be factored down, so that the combination of both responses is the same as the reduction that would have occurred due to loss.

Specifically, on a loss while in CWR state following an ECN reduction, for an implementation that uses Reno response, it would be possible to use a decrease factor of 1/(2-alpha), which would compound with the previous decrease factor of (1-alpha/2) to result in a factor of: (1 - alpha/2) / (2-alpha)) = 1/2. In integer arithmetic, this division would be possible but relatively expensive. A less expensive alternative would be a decrease factor of (2+alpha)/4, which approximates to a compounded decrease factor of 1/2 for typical low values of alpha, even up to 30%. The compound decrease factor is never greater than 1/2 and in the worst case, if alpha were 100%, it is 3/8.

If an implementation uses a CUBIC response on loss after an ECN reduction in the same round trip, in response to both ECN and loss it would not be appropriate to always aim for a combined reduction to 70%. This is because Prague's response to ECN alone can reduce cwnd to as little as 50%, so aiming for 70% would perversely require cwnd to increase on a loss in the same round as ECN marking. Experimentation is needed, but an initial proposal would be a multiplicative decrease factor of (2+alpha)/3. This would never result in an increase on loss. It would result in a combined reduction factor of about 2/3 (i.e. almost 70%) if alpha was low, rising to a combined reduction factor of 1/2 as alpha tends towards 100%.

3.2.2. ECN with Delay

Section 3.1.2 described the plans to shift between using ECN when close to the operating point and using delay by injecting paced chirps to find a new operating after the ECN signal goes silent for a few rounds. Paced chirping shifts more slowly to the new operating point the more noise there is in the delay measurements. Work is ongoing on treating any ECN marking as a complementary metric. The resulting less noisy combined metric should then allow the controller to shift more rapidly to each new operating point.

An alternative would be to combine ECN with the BBR approach, which induces a much less noisy delay signal by using less frequent but more pronounced delay spikes. The approach currently being taken is to adapt the chirp length to the degree of noise, so the chirps only become longer and/or more pronounced when necessary, for instance when faced with a discontinuous link technology such as WiFi. With multiple chirps per round, the noise can still be filtered out by averaging over them all, rather than trying to remove noise from each spike. This keeps the 'self-harm' to the minimum necessary, and ensures that capacity is always being sampled, which removes the risk of going stale.

3.3. Fall-Back on Classic ECN

The implementation of the TCP Prague CCA in Linux includes an algorithm to detect a Classic ECN AQM and fall back to Reno as a result, as required by the 'Coexistence with Classic ECN' aspect of the Prague Req 4.3. [RFC9331].

The algorithm currently used (v2) is relatively simple, but rather than describe it here, full rationale, pseudocode and explanation can be found in the technical report about it [ecn-fallback]. This also includes a selection of the evaluation results and a link to visualizations of the full results online. The current algorithm nearly always detects a Classic ECN AQM, and in the majority of the wide range of scenarios tested it is good at detecting an L4S AQM. However, it wrongly identifies an L4S AQM as Classic in a significant minority of cases when the link rate is low, or the RTT is high. The report gives ideas on how to improve detection in these scenarios, but in the mean time the algorithm has been disabled by default.

Recently, the report cited above has been updated to include new ideas on other ways to distinguish Classic from L4S AQMs. The interested reader can access it themselves, so this living document will not be further summarized here.

3.4. Further Reduced RTT-Dependence

The algorithm to reduce RTT dependence is only relevant for long-running flows. So in the current TCP Prague implementation it remains disabled for a certain number of round trips after the start of a flow, as explained in Section 2.4.4. Instead, it would be possible to make rtt_virt gradually move from the actual RTT to the target virtual RTT, or perhaps depend on other parameters of the flow. Nonetheless, just switching in the algorithm after a number of rounds works well enough. It is planned to also disable the algorithm for a similar duration if a flow becomes idle then restarts, but this is yet to be evaluated.

Prague Req 4.3. in [RFC9331]) only requires reduced RTT bias "in the range between the minimum likely RTT and typical RTTs expected in the intended deployment scenario". Nonetheless, in future it would be preferable to be able to reduce the RTT bias for high RTT flows as well.

If a step AQM is used, the congestion episodes of flows with different RTTs tend to synchronize, which exacerbates RTT bias. To prevent this two candidate approaches will need to be investigated: i) It might be sufficient to deprecate step AQMs for L4S (they are not the preferred recommendation in [RFC9332]); or ii) the virtual RTT approach of Section 2.4.4 might be usable for higher than typical RTTs as well as lower. In this latter case, (srtt/rtt_virt)^2 segments would need to be added to the window per actual RTT. The current TCP Prague implementation does not support this faster AI for RTTs longer than RTT_VIRT_MIN (25ms), due to the expected (but unverified) impact on latency overshoot and responsiveness.

3.5. Scaling Down to Fractional Windows

A modification to v5.0 of the Linux TCP stack that scales down to sub-packet windows is available for research purposes via the L4S landing page [L4S-home]. The L4S Prague Requirements in section 4.3 of [RFC9331] recommend but no longer mandate scaling down to sub-packet windows. This is because becoming unresponsive at a minimum window is a tradeoff between protecting against other unresponsive flows and the extra queue you induce by becoming unresponsive yourself. So this code is not maintained as part of the Linux implementation of TCP Prague.

Firstly, the stack has to be modifed to maintain a fractional congestion window. Then, because the ACK clock cannot work below 1 packet per RTT, the code sets the time to send each packet, then readjusts the timing as each ACK arrives (otherwise any queuing accumulates a burst in subsequent rounds). Also, additive increase of one segment does not scale below a 1-segment window. So instead of a constant additive increase, the code uses a logarithmically scaled additive increase that slowly adapts the additive increase constant to the slow start threshold. Despite these quite radical changes, the diff is surprisingly small. The design and implementation is explained in [Ahmed19], which also includes evaluation results.

4. IANA Considerations

This specification contains no IANA considerations.

5. Security Considerations

Section 3.5 on scaling down to fractional windows discusses the tradeoff in becoming unresponsive at a minium window, which causes a queue to build (harm to self and to others) but protects oneself against other unresponsive flows (whether malicious or accidental).

This draft inherits the security considerations discussed in [RFC9331] and in the L4S architecture [RFC9330]. In particular, the self-interest incentive to be responsive and minimize queuing delay, and protections against those interested in disrupting the low queuing delay of others.

6. References

6.1. Normative References

Briscoe, B., Kühlewind, M., and R. Scheffenegger, "More Accurate Explicit Congestion Notification (ECN) Feedback in TCP", Work in Progress, Internet-Draft, draft-ietf-tcpm-accurate-ecn-26, , <>.
Bradner, S., "Key words for use in RFCs to Indicate Requirement Levels", BCP 14, RFC 2119, DOI 10.17487/RFC2119, , <>.
Ramakrishnan, K., Floyd, S., and D. Black, "The Addition of Explicit Congestion Notification (ECN) to IP", RFC 3168, DOI 10.17487/RFC3168, , <>.
Black, D., "Relaxing Restrictions on Explicit Congestion Notification (ECN) Experimentation", RFC 8311, DOI 10.17487/RFC8311, , <>.
De Schepper, K. and B. Briscoe, Ed., "The Explicit Congestion Notification (ECN) Protocol for Low Latency, Low Loss, and Scalable Throughput (L4S)", RFC 9331, DOI 10.17487/RFC9331, , <>.

6.2. Informative References

Ahmed, A.S., "Extending TCP for Low Round Trip Delay", Masters Thesis, Uni Oslo , , <>.
Misund, J. and B. Briscoe, "Disentangling Flaws in Linux DCTCP", arXiv Technical Report 211.07581 [cs.NI], , <>.
Briscoe, B. and A.S. Ahmed, "TCP Prague Fall-back on Detection of a Classic ECN AQM", Technical Report TR-BB-2019-002, , <>.
Bagnulo, M. and B. Briscoe, "ECN++: Adding Explicit Congestion Notification (ECN) to TCP Control Packets", Work in Progress, Internet-Draft, draft-ietf-tcpm-generalized-ecn-12, , <>.
"L4S: Ultra-Low Queuing Delay for All", <>.
"Linux kernel tree with L4S patches", <>.
Misund, J. and B. Briscoe, "Paced Chirping - Rethinking TCP start-up", Proc. Linux Netdev 0x13 , , <>.
Shewmaker, A. G., "tcp: allow dctcp alpha to drop to zero", Linux GitHub patch; Commit: c80dbe0, , <>.
De Schepper, K., "tcp: Ensure DCTCP reacts to losses", Linux GitHub patch; Commit: aecfde2, , <>.
Briscoe, B., "Improving DCTCP/Prague Congestion Control Responsiveness", Technical Report TR-BB-2020-002, , <>.
Briscoe, B., De Schepper, K., Albisser, O., Misund, J., Tilmans, O., Kühlewind, M., and A.S. Ahmed, "Implementing the `TCP Prague' Requirements for Low Latency Low Loss Scalable Throughput (L4S)", Proc. Linux Netdev 0x13 , , <>.
Floyd, S., "HighSpeed TCP for Large Congestion Windows", RFC 3649, DOI 10.17487/RFC3649, , <>.
Kohler, E., Handley, M., and S. Floyd, "Datagram Congestion Control Protocol (DCCP)", RFC 4340, DOI 10.17487/RFC4340, , <>.
Floyd, S. and M. Allman, "Specifying New Congestion Control Algorithms", BCP 133, RFC 5033, DOI 10.17487/RFC5033, , <>.
Allman, M., Paxson, V., and E. Blanton, "TCP Congestion Control", RFC 5681, DOI 10.17487/RFC5681, , <>.
Westerlund, M., Johansson, I., Perkins, C., O'Hanlon, P., and K. Carlberg, "Explicit Congestion Notification (ECN) for RTP over UDP", RFC 6679, DOI 10.17487/RFC6679, , <>.
Chu, J., Dukkipati, N., Cheng, Y., and M. Mathis, "Increasing TCP's Initial Window", RFC 6928, DOI 10.17487/RFC6928, , <>.
Leiba, B., "Ambiguity of Uppercase vs Lowercase in RFC 2119 Key Words", BCP 14, RFC 8174, DOI 10.17487/RFC8174, , <>.
Bensley, S., Thaler, D., Balasubramanian, P., Eggert, L., and G. Judd, "Data Center TCP (DCTCP): TCP Congestion Control for Data Centers", RFC 8257, DOI 10.17487/RFC8257, , <>.
Johansson, I. and Z. Sarker, "Self-Clocked Rate Adaptation for Multimedia", RFC 8298, DOI 10.17487/RFC8298, , <>.
Sarker, Z., Perkins, C., Singh, V., and M. Ramalho, "RTP Control Protocol (RTCP) Feedback for Congestion Control", RFC 8888, DOI 10.17487/RFC8888, , <>.
Cheng, Y., Cardwell, N., Dukkipati, N., and P. Jha, "The RACK-TLP Loss Detection Algorithm for TCP", RFC 8985, DOI 10.17487/RFC8985, , <>.
Iyengar, J., Ed. and M. Thomson, Ed., "QUIC: A UDP-Based Multiplexed and Secure Transport", RFC 9000, DOI 10.17487/RFC9000, , <>.
Stewart, R., Tüxen, M., and K. Nielsen, "Stream Control Transmission Protocol", RFC 9260, DOI 10.17487/RFC9260, , <>.
Briscoe, B., Ed., De Schepper, K., Bagnulo, M., and G. White, "Low Latency, Low Loss, and Scalable Throughput (L4S) Internet Service: Architecture", RFC 9330, DOI 10.17487/RFC9330, , <>.
De Schepper, K., Briscoe, B., Ed., and G. White, "Dual-Queue Coupled Active Queue Management (AQM) for Low Latency, Low Loss, and Scalable Throughput (L4S)", RFC 9332, DOI 10.17487/RFC9332, , <>.
Balasubramanian, P., Huang, Y., and M. Olson, "HyStart++: Modified Slow Start for TCP", RFC 9406, DOI 10.17487/RFC9406, , <>.
Xu, L., Ha, S., Rhee, I., Goel, V., and L. Eggert, Ed., "CUBIC for Fast and Long-Distance Networks", RFC 9438, DOI 10.17487/RFC9438, , <>.
Briscoe, B. and K. De Schepper, "Resolving Tensions between Congestion Control Scaling Requirements", Simula Technical Report TR-CS-2016-001; arXiv:1904.07605, , <>.

Comments and Contributions Solicited

This section is to be removed before publishing as an RFC.

Comments and questions are encouraged and very welcome. They can be addressed to the IRTF Internet Congestion Control Research Group's mailing list <>, and/or to the authors via <>.

Contributions of design ideas and/or code are also encouraged and welcome. During the drafting process, the intention is to gather experience into this document from a wider set of Prague congestion control implementations.


Thanks to Neal Cardwell for review comments.

The following contributed implementations and evaluations that validated and helped to improve this specification:

  • Olivier Tilmans <> of Nokia Bell Labs, Belgium, prepared and maintains the Linux implementation of TCP Prague.

  • Koen De Schepper <> of Nokia Bell Labs, Belgium, contributed to the Linux implementation of TCP Prague.

  • Joakim Misund <> of Uni Oslo, Norway, wrote the Linux paced chirping code.

  • Asad Sajjad Ahmed <>, Independent, Norway, wrote the Linux code that maintains a sub-packet window.

  • Vidhi Goel <> of Apple Inc, Cupertino, wrote and maintains the Apple implmentation of QUIC Prague.

Bob Briscoe's contribution was part-funded by the Comcast Innovation Fund and part-funded by Apple Inc. The views expressed here are solely those of the authors.

Authors' Addresses

Koen De Schepper
Nokia Bell Labs
Olivier Tilmans
Nokia Bell Labs
Bob Briscoe (editor)
United Kingdom
Vidhi Goel
Apple Inc
United States