diff options
Diffstat (limited to 'Documentation')
| -rw-r--r-- | Documentation/RCU/RTFP.txt | 149 | ||||
| -rw-r--r-- | Documentation/RCU/checklist.txt | 18 | ||||
| -rw-r--r-- | Documentation/devicetree/bindings/net/micrel-ks8851.txt | 1 | ||||
| -rw-r--r-- | Documentation/kernel-per-CPU-kthreads.txt | 13 | ||||
| -rw-r--r-- | Documentation/memory-barriers.txt | 137 | ||||
| -rw-r--r-- | Documentation/networking/netlink_mmap.txt | 4 | ||||
| -rw-r--r-- | Documentation/networking/packet_mmap.txt | 2 | ||||
| -rw-r--r-- | Documentation/networking/timestamping.txt | 52 | ||||
| -rw-r--r-- | Documentation/sysctl/kernel.txt | 15 |
9 files changed, 289 insertions, 102 deletions
diff --git a/Documentation/RCU/RTFP.txt b/Documentation/RCU/RTFP.txt index 273e654d7d08..2f0fcb2112d2 100644 --- a/Documentation/RCU/RTFP.txt +++ b/Documentation/RCU/RTFP.txt @@ -31,6 +31,14 @@ has lapsed, so this approach may be used in non-GPL software, if desired. (In contrast, implementation of RCU is permitted only in software licensed under either GPL or LGPL. Sorry!!!) +In 1987, Rashid et al. described lazy TLB-flush [RichardRashid87a]. +At first glance, this has nothing to do with RCU, but nevertheless +this paper helped inspire the update-side batching used in the later +RCU implementation in DYNIX/ptx. In 1988, Barbara Liskov published +a description of Argus that noted that use of out-of-date values can +be tolerated in some situations. Thus, this paper provides some early +theoretical justification for use of stale data. + In 1990, Pugh [Pugh90] noted that explicitly tracking which threads were reading a given data structure permitted deferred free to operate in the presence of non-terminating threads. However, this explicit @@ -41,11 +49,11 @@ providing a fine-grained locking design, however, it would be interesting to see how much of the performance advantage reported in 1990 remains today. -At about this same time, Adams [Adams91] described ``chaotic relaxation'', -where the normal barriers between successive iterations of convergent -numerical algorithms are relaxed, so that iteration $n$ might use -data from iteration $n-1$ or even $n-2$. This introduces error, -which typically slows convergence and thus increases the number of +At about this same time, Andrews [Andrews91textbook] described ``chaotic +relaxation'', where the normal barriers between successive iterations +of convergent numerical algorithms are relaxed, so that iteration $n$ +might use data from iteration $n-1$ or even $n-2$. This introduces +error, which typically slows convergence and thus increases the number of iterations required. However, this increase is sometimes more than made up for by a reduction in the number of expensive barrier operations, which are otherwise required to synchronize the threads at the end @@ -55,7 +63,8 @@ is thus inapplicable to most data structures in operating-system kernels. In 1992, Henry (now Alexia) Massalin completed a dissertation advising parallel programmers to defer processing when feasible to simplify -synchronization. RCU makes extremely heavy use of this advice. +synchronization [HMassalinPhD]. RCU makes extremely heavy use of +this advice. In 1993, Jacobson [Jacobson93] verbally described what is perhaps the simplest deferred-free technique: simply waiting a fixed amount of time @@ -90,27 +99,29 @@ mechanism, which is quite similar to RCU [Gamsa99]. These operating systems made pervasive use of RCU in place of "existence locks", which greatly simplifies locking hierarchies and helps avoid deadlocks. -2001 saw the first RCU presentation involving Linux [McKenney01a] -at OLS. The resulting abundance of RCU patches was presented the -following year [McKenney02a], and use of RCU in dcache was first -described that same year [Linder02a]. +The year 2000 saw an email exchange that would likely have +led to yet another independent invention of something like RCU +[RustyRussell2000a,RustyRussell2000b]. Instead, 2001 saw the first +RCU presentation involving Linux [McKenney01a] at OLS. The resulting +abundance of RCU patches was presented the following year [McKenney02a], +and use of RCU in dcache was first described that same year [Linder02a]. Also in 2002, Michael [Michael02b,Michael02a] presented "hazard-pointer" techniques that defer the destruction of data structures to simplify non-blocking synchronization (wait-free synchronization, lock-free synchronization, and obstruction-free synchronization are all examples of -non-blocking synchronization). In particular, this technique eliminates -locking, reduces contention, reduces memory latency for readers, and -parallelizes pipeline stalls and memory latency for writers. However, -these techniques still impose significant read-side overhead in the -form of memory barriers. Researchers at Sun worked along similar lines -in the same timeframe [HerlihyLM02]. These techniques can be thought -of as inside-out reference counts, where the count is represented by the -number of hazard pointers referencing a given data structure rather than -the more conventional counter field within the data structure itself. -The key advantage of inside-out reference counts is that they can be -stored in immortal variables, thus allowing races between access and -deletion to be avoided. +non-blocking synchronization). The corresponding journal article appeared +in 2004 [MagedMichael04a]. This technique eliminates locking, reduces +contention, reduces memory latency for readers, and parallelizes pipeline +stalls and memory latency for writers. However, these techniques still +impose significant read-side overhead in the form of memory barriers. +Researchers at Sun worked along similar lines in the same timeframe +[HerlihyLM02]. These techniques can be thought of as inside-out reference +counts, where the count is represented by the number of hazard pointers +referencing a given data structure rather than the more conventional +counter field within the data structure itself. The key advantage +of inside-out reference counts is that they can be stored in immortal +variables, thus allowing races between access and deletion to be avoided. By the same token, RCU can be thought of as a "bulk reference count", where some form of reference counter covers all reference by a given CPU @@ -123,8 +134,10 @@ can be thought of in other terms as well. In 2003, the K42 group described how RCU could be used to create hot-pluggable implementations of operating-system functions [Appavoo03a]. -Later that year saw a paper describing an RCU implementation of System -V IPC [Arcangeli03], and an introduction to RCU in Linux Journal +Later that year saw a paper describing an RCU implementation +of System V IPC [Arcangeli03] (following up on a suggestion by +Hugh Dickins [Dickins02a] and an implementation by Mingming Cao +[MingmingCao2002IPCRCU]), and an introduction to RCU in Linux Journal [McKenney03a]. 2004 has seen a Linux-Journal article on use of RCU in dcache @@ -383,6 +396,21 @@ for Programming Languages and Operating Systems}" } } +@phdthesis{HMassalinPhD +,author="H. Massalin" +,title="Synthesis: An Efficient Implementation of Fundamental Operating +System Services" +,school="Columbia University" +,address="New York, NY" +,year="1992" +,annotation={ + Mondo optimizing compiler. + Wait-free stuff. + Good advice: defer work to avoid synchronization. See page 90 + (PDF page 106), Section 5.4, fourth bullet point. +} +} + @unpublished{Jacobson93 ,author="Van Jacobson" ,title="Avoid Read-Side Locking Via Delayed Free" @@ -671,6 +699,20 @@ Orran Krieger and Rusty Russell and Dipankar Sarma and Maneesh Soni" [Viewed October 18, 2004]" } +@conference{Michael02b +,author="Maged M. Michael" +,title="High Performance Dynamic Lock-Free Hash Tables and List-Based Sets" +,Year="2002" +,Month="August" +,booktitle="{Proceedings of the 14\textsuperscript{th} Annual ACM +Symposium on Parallel +Algorithms and Architecture}" +,pages="73-82" +,annotation={ +Like the title says... +} +} + @Conference{Linder02a ,Author="Hanna Linder and Dipankar Sarma and Maneesh Soni" ,Title="Scalability of the Directory Entry Cache" @@ -727,6 +769,24 @@ Andrea Arcangeli and Andi Kleen and Orran Krieger and Rusty Russell" } } +@conference{Michael02a +,author="Maged M. Michael" +,title="Safe Memory Reclamation for Dynamic Lock-Free Objects Using Atomic +Reads and Writes" +,Year="2002" +,Month="August" +,booktitle="{Proceedings of the 21\textsuperscript{st} Annual ACM +Symposium on Principles of Distributed Computing}" +,pages="21-30" +,annotation={ + Each thread keeps an array of pointers to items that it is + currently referencing. Sort of an inside-out garbage collection + mechanism, but one that requires the accessing code to explicitly + state its needs. Also requires read-side memory barriers on + most architectures. +} +} + @unpublished{Dickins02a ,author="Hugh Dickins" ,title="Use RCU for System-V IPC" @@ -735,6 +795,17 @@ Andrea Arcangeli and Andi Kleen and Orran Krieger and Rusty Russell" ,note="private communication" } +@InProceedings{HerlihyLM02 +,author={Maurice Herlihy and Victor Luchangco and Mark Moir} +,title="The Repeat Offender Problem: A Mechanism for Supporting Dynamic-Sized, +Lock-Free Data Structures" +,booktitle={Proceedings of 16\textsuperscript{th} International +Symposium on Distributed Computing} +,year=2002 +,month="October" +,pages="339-353" +} + @unpublished{Sarma02b ,Author="Dipankar Sarma" ,Title="Some dcache\_rcu benchmark numbers" @@ -749,6 +820,19 @@ Andrea Arcangeli and Andi Kleen and Orran Krieger and Rusty Russell" } } +@unpublished{MingmingCao2002IPCRCU +,Author="Mingming Cao" +,Title="[PATCH]updated ipc lock patch" +,month="October" +,year="2002" +,note="Available: +\url{https://lkml.org/lkml/2002/10/24/262} +[Viewed February 15, 2014]" +,annotation={ + Mingming Cao's patch to introduce RCU to SysV IPC. +} +} + @unpublished{LinusTorvalds2003a ,Author="Linus Torvalds" ,Title="Re: {[PATCH]} small fixes in brlock.h" @@ -982,6 +1066,23 @@ Realtime Applications" } } +@article{MagedMichael04a +,author="Maged M. Michael" +,title="Hazard Pointers: Safe Memory Reclamation for Lock-Free Objects" +,Year="2004" +,Month="June" +,journal="IEEE Transactions on Parallel and Distributed Systems" +,volume="15" +,number="6" +,pages="491-504" +,url="Available: +\url{http://www.research.ibm.com/people/m/michael/ieeetpds-2004.pdf} +[Viewed March 1, 2005]" +,annotation={ + New canonical hazard-pointer citation. +} +} + @phdthesis{PaulEdwardMcKenneyPhD ,author="Paul E. McKenney" ,title="Exploiting Deferred Destruction: diff --git a/Documentation/RCU/checklist.txt b/Documentation/RCU/checklist.txt index 91266193b8f4..9d10d1db16a5 100644 --- a/Documentation/RCU/checklist.txt +++ b/Documentation/RCU/checklist.txt @@ -256,10 +256,10 @@ over a rather long period of time, but improvements are always welcome! variations on this theme. b. Limiting update rate. For example, if updates occur only - once per hour, then no explicit rate limiting is required, - unless your system is already badly broken. The dcache - subsystem takes this approach -- updates are guarded - by a global lock, limiting their rate. + once per hour, then no explicit rate limiting is + required, unless your system is already badly broken. + Older versions of the dcache subsystem take this approach, + guarding updates with a global lock, limiting their rate. c. Trusted update -- if updates can only be done manually by superuser or some other trusted user, then it might not @@ -268,7 +268,8 @@ over a rather long period of time, but improvements are always welcome! the machine. d. Use call_rcu_bh() rather than call_rcu(), in order to take - advantage of call_rcu_bh()'s faster grace periods. + advantage of call_rcu_bh()'s faster grace periods. (This + is only a partial solution, though.) e. Periodically invoke synchronize_rcu(), permitting a limited number of updates per grace period. @@ -276,6 +277,13 @@ over a rather long period of time, but improvements are always welcome! The same cautions apply to call_rcu_bh(), call_rcu_sched(), call_srcu(), and kfree_rcu(). + Note that although these primitives do take action to avoid memory + exhaustion when any given CPU has too many callbacks, a determined + user could still exhaust memory. This is especially the case + if a system with a large number of CPUs has been configured to + offload all of its RCU callbacks onto a single CPU, or if the + system has relatively little free memory. + 9. All RCU list-traversal primitives, which include rcu_dereference(), list_for_each_entry_rcu(), and list_for_each_safe_rcu(), must be either within an RCU read-side diff --git a/Documentation/devicetree/bindings/net/micrel-ks8851.txt b/Documentation/devicetree/bindings/net/micrel-ks8851.txt index 11ace3c3d805..4fc392763611 100644 --- a/Documentation/devicetree/bindings/net/micrel-ks8851.txt +++ b/Documentation/devicetree/bindings/net/micrel-ks8851.txt @@ -7,3 +7,4 @@ Required properties: Optional properties: - local-mac-address : Ethernet mac address to use +- vdd-supply: supply for Ethernet mac diff --git a/Documentation/kernel-per-CPU-kthreads.txt b/Documentation/kernel-per-CPU-kthreads.txt index 827104fb9364..f3cd299fcc41 100644 --- a/Documentation/kernel-per-CPU-kthreads.txt +++ b/Documentation/kernel-per-CPU-kthreads.txt @@ -162,7 +162,18 @@ Purpose: Execute workqueue requests To reduce its OS jitter, do any of the following: 1. Run your workload at a real-time priority, which will allow preempting the kworker daemons. -2. Do any of the following needed to avoid jitter that your +2. A given workqueue can be made visible in the sysfs filesystem + by passing the WQ_SYSFS to that workqueue's alloc_workqueue(). + Such a workqueue can be confined to a given subset of the + CPUs using the /sys/devices/virtual/workqueue/*/cpumask sysfs + files. The set of WQ_SYSFS workqueues can be displayed using + "ls sys/devices/virtual/workqueue". That said, the workqueues + maintainer would like to caution people against indiscriminately + sprinkling WQ_SYSFS across all the workqueues. The reason for + caution is that it is easy to add WQ_SYSFS, but because sysfs is + part of the formal user/kernel API, it can be nearly impossible + to remove it, even if its addition was a mistake. +3. Do any of the following needed to avoid jitter that your application cannot tolerate: a. Build your kernel with CONFIG_SLUB=y rather than CONFIG_SLAB=y, thus avoiding the slab allocator's periodic diff --git a/Documentation/memory-barriers.txt b/Documentation/memory-barriers.txt index 102dc19c4119..11c1d2049662 100644 --- a/Documentation/memory-barriers.txt +++ b/Documentation/memory-barriers.txt @@ -608,26 +608,30 @@ as follows: b = p; /* BUG: Compiler can reorder!!! */ do_something(); -The solution is again ACCESS_ONCE(), which preserves the ordering between -the load from variable 'a' and the store to variable 'b': +The solution is again ACCESS_ONCE() and barrier(), which preserves the +ordering between the load from variable 'a' and the store to variable 'b': q = ACCESS_ONCE(a); if (q) { + barrier(); ACCESS_ONCE(b) = p; do_something(); } else { + barrier(); ACCESS_ONCE(b) = p; do_something_else(); } -You could also use barrier() to prevent the compiler from moving -the stores to variable 'b', but barrier() would not prevent the -compiler from proving to itself that a==1 always, so ACCESS_ONCE() -is also needed. +The initial ACCESS_ONCE() is required to prevent the compiler from +proving the value of 'a', and the pair of barrier() invocations are +required to prevent the compiler from pulling the two identical stores +to 'b' out from the legs of the "if" statement. It is important to note that control dependencies absolutely require a a conditional. For example, the following "optimized" version of -the above example breaks ordering: +the above example breaks ordering, which is why the barrier() invocations +are absolutely required if you have identical stores in both legs of +the "if" statement: q = ACCESS_ONCE(a); ACCESS_ONCE(b) = p; /* BUG: No ordering vs. load from a!!! */ @@ -643,9 +647,11 @@ It is of course legal for the prior load to be part of the conditional, for example, as follows: if (ACCESS_ONCE(a) > 0) { + barrier(); ACCESS_ONCE(b) = q / 2; do_something(); } else { + barrier(); ACCESS_ONCE(b) = q / 3; do_something_else(); } @@ -659,9 +665,11 @@ the needed conditional. For example: q = ACCESS_ONCE(a); if (q % MAX) { + barrier(); ACCESS_ONCE(b) = p; do_something(); } else { + barrier(); ACCESS_ONCE(b) = p; do_something_else(); } @@ -723,8 +731,13 @@ In summary: use smb_rmb(), smp_wmb(), or, in the case of prior stores and later loads, smp_mb(). + (*) If both legs of the "if" statement begin with identical stores + to the same variable, a barrier() statement is required at the + beginning of each leg of the "if" statement. + (*) Control dependencies require at least one run-time conditional - between the prior load and the subsequent store. If the compiler + between the prior load and the subsequent store, and this + conditional must involve the prior load. If the compiler is able to optimize the conditional away, it will have also optimized away the ordering. Careful use of ACCESS_ONCE() can help to preserve the needed conditional. @@ -1249,6 +1262,23 @@ The ACCESS_ONCE() function can prevent any number of optimizations that, while perfectly safe in single-threaded code, can be fatal in concurrent code. Here are some examples of these sorts of optimizations: + (*) The compiler is within its rights to reorder loads and stores + to the same variable, and in some cases, the CPU is within its + rights to reorder loads to the same variable. This means that + the following code: + + a[0] = x; + a[1] = x; + + Might result in an older value of x stored in a[1] than in a[0]. + Prevent both the compiler and the CPU from doing this as follows: + + a[0] = ACCESS_ONCE(x); + a[1] = ACCESS_ONCE(x); + + In short, ACCESS_ONCE() provides cache coherence for accesses from + multiple CPUs to a single variable. + (*) The compiler is within its rights to merge successive loads from the same variable. Such merging can cause the compiler to "optimize" the following code: @@ -1644,12 +1674,12 @@ for each construct. These operations all imply certain barriers: Memory operations issued after the ACQUIRE will be completed after the ACQUIRE operation has completed. - Memory operations issued before the ACQUIRE may be completed after the - ACQUIRE operation has completed. An smp_mb__before_spinlock(), combined - with a following ACQUIRE, orders prior loads against subsequent stores and - stores and prior stores against subsequent stores. Note that this is - weaker than smp_mb()! The smp_mb__before_spinlock() primitive is free on - many architectures. + Memory operations issued before the ACQUIRE may be completed after + the ACQUIRE operation has completed. An smp_mb__before_spinlock(), + combined with a following ACQUIRE, orders prior loads against + subsequent loads and stores and also orders prior stores against + subsequent stores. Note that this is weaker than smp_mb()! The + smp_mb__before_spinlock() primitive is free on many architectures. (2) RELEASE operation implication: @@ -1694,24 +1724,21 @@ may occur as: ACQUIRE M, STORE *B, STORE *A, RELEASE M -This same reordering can of course occur if the lock's ACQUIRE and RELEASE are -to the same lock variable, but only from the perspective of another CPU not -holding that lock. - -In short, a RELEASE followed by an ACQUIRE may -not- be assumed to be a full -memory barrier because it is possible for a preceding RELEASE to pass a -later ACQUIRE from the viewpoint of the CPU, but not from the viewpoint -of the compiler. Note that deadlocks cannot be introduced by this -interchange because if such a deadlock threatened, the RELEASE would -simply complete. - -If it is necessary for a RELEASE-ACQUIRE pair to produce a full barrier, the -ACQUIRE can be followed by an smp_mb__after_unlock_lock() invocation. This -will produce a full barrier if either (a) the RELEASE and the ACQUIRE are -executed by the same CPU or task, or (b) the RELEASE and ACQUIRE act on the -same variable. The smp_mb__after_unlock_lock() primitive is free on many -architectures. Without smp_mb__after_unlock_lock(), the critical sections -corresponding to the RELEASE and the ACQUIRE can cross: +When the ACQUIRE and RELEASE are a lock acquisition and release, +respectively, this same reordering can occur if the lock's ACQUIRE and +RELEASE are to the same lock variable, but only from the perspective of +another CPU not holding that lock. In short, a ACQUIRE followed by an +RELEASE may -not- be assumed to be a full memory barrier. + +Similarly, the reverse case of a RELEASE followed by an ACQUIRE does not +imply a full memory barrier. If it is necessary for a RELEASE-ACQUIRE +pair to produce a full barrier, the ACQUIRE can be followed by an +smp_mb__after_unlock_lock() invocation. This will produce a full barrier +if either (a) the RELEASE and the ACQUIRE are executed by the same +CPU or task, or (b) the RELEASE and ACQUIRE act on the same variable. +The smp_mb__after_unlock_lock() primitive is free on many architectures. +Without smp_mb__after_unlock_lock(), the CPU's execution of the critical +sections corresponding to the RELEASE and the ACQUIRE can cross, so that: *A = a; RELEASE M @@ -1722,7 +1749,36 @@ could occur as: ACQUIRE N, STORE *B, STORE *A, RELEASE M -With smp_mb__after_unlock_lock(), they cannot, so that: +It might appear that this reordering could introduce a deadlock. +However, this cannot happen because if such a deadlock threatened, +the RELEASE would simply complete, thereby avoiding the deadlock. + + Why does this work? + + One key point is that we are only talking about the CPU doing + the reordering, not the compiler. If the compiler (or, for + that matter, the developer) switched the operations, deadlock + -could- occur. + + But suppose the CPU reordered the operations. In this case, + the unlock precedes the lock in the assembly code. The CPU + simply elected to try executing the later lock operation first. + If there is a deadlock, this lock operation will simply spin (or + try to sleep, but more on that later). The CPU will eventually + execute the unlock operation (which preceded the lock operation + in the assembly code), which will unravel the potential deadlock, + allowing the lock operation to succeed. + + But what if the lock is a sleeplock? In that case, the code will + try to enter the scheduler, where it will eventually encounter + a memory barrier, which will force the earlier unlock operation + to complete, again unraveling the deadlock. There might be + a sleep-unlock race, but the locking primitive needs to resolve + such races properly in any case. + +With smp_mb__after_unlock_lock(), the two critical sections cannot overlap. +For example, with the following code, the store to *A will always be +seen by other CPUs before the store to *B: *A = a; RELEASE M @@ -1730,13 +1786,18 @@ With smp_mb__after_unlock_lock(), they cannot, so that: smp_mb__after_unlock_lock(); *B = b; -will always occur as either of the following: +The operations will always occur in one of the following orders: - STORE *A, RELEASE, ACQUIRE, STORE *B - STORE *A, ACQUIRE, RELEASE, STORE *B + STORE *A, RELEASE, ACQUIRE, smp_mb__after_unlock_lock(), STORE *B + STORE *A, ACQUIRE, RELEASE, smp_mb__after_unlock_lock(), STORE *B + ACQUIRE, STORE *A, RELEASE, smp_mb__after_unlock_lock(), STORE *B If the RELEASE and ACQUIRE were instead both operating on the same lock -variable, only the first of these two alternatives can occur. +variable, only the first of these alternatives can occur. In addition, +the more strongly ordered systems may rule out some of the above orders. +But in any case, as noted earlier, the smp_mb__after_unlock_lock() +ensures that the store to *A will always be seen as happening before +the store to *B. Locks and semaphores may not provide any guarantee of ordering on UP compiled systems, and so cannot be counted on in such a situation to actually achieve @@ -2757,7 +2818,7 @@ in that order, but, without intervention, the sequence may have almost any combination of elements combined or discarded, provided the program's view of the world remains consistent. Note that ACCESS_ONCE() is -not- optional in the above example, as there are architectures where a given CPU might -interchange successive loads to the same location. On such architectures, +reorder successive loads to the same location. On such architectures, ACCESS_ONCE() does whatever is necessary to prevent this, for example, on Itanium the volatile casts used by ACCESS_ONCE() cause GCC to emit the special ld.acq and st.rel instructions that prevent such reordering. diff --git a/Documentation/networking/netlink_mmap.txt b/Documentation/networking/netlink_mmap.txt index b26122973525..c6af4bac5aa8 100644 --- a/Documentation/networking/netlink_mmap.txt +++ b/Documentation/networking/netlink_mmap.txt @@ -226,9 +226,9 @@ Ring setup: void *rx_ring, *tx_ring; /* Configure ring parameters */ - if (setsockopt(fd, NETLINK_RX_RING, &req, sizeof(req)) < 0) + if (setsockopt(fd, SOL_NETLINK, NETLINK_RX_RING, &req, sizeof(req)) < 0) exit(1); - if (setsockopt(fd, NETLINK_TX_RING, &req, sizeof(req)) < 0) + if (setsockopt(fd, SOL_NETLINK, NETLINK_TX_RING, &req, sizeof(req)) < 0) exit(1) /* Calculate size of each individual ring */ diff --git a/Documentation/networking/packet_mmap.txt b/Documentation/networking/packet_mmap.txt index 1404674c0a02..6fea79efb4cb 100644 --- a/Documentation/networking/packet_mmap.txt +++ b/Documentation/networking/packet_mmap.txt @@ -453,7 +453,7 @@ TP_STATUS_COPY : This flag indicates that the frame (and associated enabled previously with setsockopt() and the PACKET_COPY_THRESH option. - The number of frames than can be buffered to + The number of frames that can be buffered to be read with recvfrom is limited like a normal socket. See the SO_RCVBUF option in the socket (7) man page. diff --git a/Documentation/networking/timestamping.txt b/Documentation/networking/timestamping.txt index 661d3c316a17..048c92b487f6 100644 --- a/Documentation/networking/timestamping.txt +++ b/Documentation/networking/timestamping.txt @@ -21,26 +21,38 @@ has such a feature). SO_TIMESTAMPING: -Instructs the socket layer which kind of information is wanted. The -parameter is an integer with some of the following bits set. Setting -other bits is an error and doesn't change the current state. - -SOF_TIMESTAMPING_TX_HARDWARE: try to obtain send time stamp in hardware -SOF_TIMESTAMPING_TX_SOFTWARE: if SOF_TIMESTAMPING_TX_HARDWARE is off or - fails, then do it in software -SOF_TIMESTAMPING_RX_HARDWARE: return the original, unmodified time stamp - as generated by the hardware -SOF_TIMESTAMPING_RX_SOFTWARE: if SOF_TIMESTAMPING_RX_HARDWARE is off or - fails, then do it in software -SOF_TIMESTAMPING_RAW_HARDWARE: return original raw hardware time stamp -SOF_TIMESTAMPING_SYS_HARDWARE: return hardware time stamp transformed to - the system time base -SOF_TIMESTAMPING_SOFTWARE: return system time stamp generated in - software - -SOF_TIMESTAMPING_TX/RX determine how time stamps are generated. -SOF_TIMESTAMPING_RAW/SYS determine how they are reported in the -following control message: +Instructs the socket layer which kind of information should be collected +and/or reported. The parameter is an integer with some of the following +bits set. Setting other bits is an error and doesn't change the current +state. + +Four of the bits are requests to the stack to try to generate +timestamps. Any combination of them is valid. + +SOF_TIMESTAMPING_TX_HARDWARE: try to obtain send time stamps in hardware +SOF_TIMESTAMPING_TX_SOFTWARE: try to obtain send time stamps in software +SOF_TIMESTAMPING_RX_HARDWARE: try to obtain receive time stamps in hardware +SOF_TIMESTAMPING_RX_SOFTWARE: try to obtain receive time stamps in software + +The other three bits control which timestamps will be reported in a +generated control message. If none of these bits are set or if none of +the set bits correspond to data that is available, then the control +message will not be generated: + +SOF_TIMESTAMPING_SOFTWARE: report systime if available +SOF_TIMESTAMPING_SYS_HARDWARE: report hwtimetrans if available +SOF_TIMESTAMPING_RAW_HARDWARE: report hwtimeraw if available + +It is worth noting that timestamps may be collected for reasons other +than being requested by a particular socket with +SOF_TIMESTAMPING_[TR]X_(HARD|SOFT)WARE. For example, most drivers that +can generate hardware receive timestamps ignore +SOF_TIMESTAMPING_RX_HARDWARE. It is still a good idea to set that flag +in case future drivers pay attention. + +If timestamps are reported, they will appear in a control message with +cmsg_level==SOL_SOCKET, cmsg_type==SO_TIMESTAMPING, and a payload like +this: struct scm_timestamping { struct timespec systime; diff --git a/Documentation/sysctl/kernel.txt b/Documentation/sysctl/kernel.txt index e55124e7c40c..ec8be46bf48d 100644 --- a/Documentation/sysctl/kernel.txt +++ b/Documentation/sysctl/kernel.txt @@ -320,10 +320,11 @@ This file shows up if CONFIG_DETECT_HUNG_TASK is enabled. ============================================================== -hung_task_warning: +hung_task_warnings: The maximum number of warnings to report. During a check interval -When this value is reached, no more the warnings will be reported. +if a hung task is detected, this value is decreased by 1. +When this value reaches 0, no more warnings will be reported. This file shows up if CONFIG_DETECT_HUNG_TASK is enabled. -1: report an infinite number of warnings. @@ -441,8 +442,7 @@ feature should be disabled. Otherwise, if the system overhead from the feature is too high then the rate the kernel samples for NUMA hinting faults may be controlled by the numa_balancing_scan_period_min_ms, numa_balancing_scan_delay_ms, numa_balancing_scan_period_max_ms, -numa_balancing_scan_size_mb, numa_balancing_settle_count sysctls and -numa_balancing_migrate_deferred. +numa_balancing_scan_size_mb, and numa_balancing_settle_count sysctls. ============================================================== @@ -483,13 +483,6 @@ rate for each task. numa_balancing_scan_size_mb is how many megabytes worth of pages are scanned for a given scan. -numa_balancing_migrate_deferred is how many page migrations get skipped -unconditionally, after a page migration is skipped because a page is shared -with other tasks. This reduces page migration overhead, and determines -how much stronger the "move task near its memory" policy scheduler becomes, -versus the "move memory near its task" memory management policy, for workloads -with shared memory. - ============================================================== osrelease, ostype & version: |