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1 ============================
2 LINUX KERNEL MEMORY BARRIERS
3 ============================
5 By: David Howells <dhowells@redhat.com>
6 Paul E. McKenney <paulmck@linux.vnet.ibm.com>
8 Contents:
10 (*) Abstract memory access model.
12 - Device operations.
13 - Guarantees.
15 (*) What are memory barriers?
17 - Varieties of memory barrier.
18 - What may not be assumed about memory barriers?
19 - Data dependency barriers.
20 - Control dependencies.
21 - SMP barrier pairing.
22 - Examples of memory barrier sequences.
23 - Read memory barriers vs load speculation.
24 - Transitivity
26 (*) Explicit kernel barriers.
28 - Compiler barrier.
29 - CPU memory barriers.
30 - MMIO write barrier.
32 (*) Implicit kernel memory barriers.
34 - Locking functions.
35 - Interrupt disabling functions.
36 - Sleep and wake-up functions.
37 - Miscellaneous functions.
39 (*) Inter-CPU locking barrier effects.
41 - Locks vs memory accesses.
42 - Locks vs I/O accesses.
44 (*) Where are memory barriers needed?
46 - Interprocessor interaction.
47 - Atomic operations.
48 - Accessing devices.
49 - Interrupts.
51 (*) Kernel I/O barrier effects.
53 (*) Assumed minimum execution ordering model.
55 (*) The effects of the cpu cache.
57 - Cache coherency.
58 - Cache coherency vs DMA.
59 - Cache coherency vs MMIO.
61 (*) The things CPUs get up to.
63 - And then there's the Alpha.
65 (*) Example uses.
67 - Circular buffers.
69 (*) References.
72 ============================
73 ABSTRACT MEMORY ACCESS MODEL
74 ============================
76 Consider the following abstract model of the system:
78 : :
79 : :
80 : :
81 +-------+ : +--------+ : +-------+
82 | | : | | : | |
83 | | : | | : | |
84 | CPU 1 |<----->| Memory |<----->| CPU 2 |
85 | | : | | : | |
86 | | : | | : | |
87 +-------+ : +--------+ : +-------+
88 ^ : ^ : ^
89 | : | : |
90 | : | : |
91 | : v : |
92 | : +--------+ : |
93 | : | | : |
94 | : | | : |
95 +---------->| Device |<----------+
96 : | | :
97 : | | :
98 : +--------+ :
99 : :
101 Each CPU executes a program that generates memory access operations. In the
102 abstract CPU, memory operation ordering is very relaxed, and a CPU may actually
103 perform the memory operations in any order it likes, provided program causality
104 appears to be maintained. Similarly, the compiler may also arrange the
105 instructions it emits in any order it likes, provided it doesn't affect the
106 apparent operation of the program.
108 So in the above diagram, the effects of the memory operations performed by a
109 CPU are perceived by the rest of the system as the operations cross the
110 interface between the CPU and rest of the system (the dotted lines).
113 For example, consider the following sequence of events:
115 CPU 1 CPU 2
116 =============== ===============
117 { A == 1; B == 2 }
118 A = 3; x = B;
119 B = 4; y = A;
121 The set of accesses as seen by the memory system in the middle can be arranged
122 in 24 different combinations:
124 STORE A=3, STORE B=4, y=LOAD A->3, x=LOAD B->4
125 STORE A=3, STORE B=4, x=LOAD B->4, y=LOAD A->3
126 STORE A=3, y=LOAD A->3, STORE B=4, x=LOAD B->4
127 STORE A=3, y=LOAD A->3, x=LOAD B->2, STORE B=4
128 STORE A=3, x=LOAD B->2, STORE B=4, y=LOAD A->3
129 STORE A=3, x=LOAD B->2, y=LOAD A->3, STORE B=4
130 STORE B=4, STORE A=3, y=LOAD A->3, x=LOAD B->4
131 STORE B=4, ...
132 ...
134 and can thus result in four different combinations of values:
136 x == 2, y == 1
137 x == 2, y == 3
138 x == 4, y == 1
139 x == 4, y == 3
142 Furthermore, the stores committed by a CPU to the memory system may not be
143 perceived by the loads made by another CPU in the same order as the stores were
144 committed.
147 As a further example, consider this sequence of events:
149 CPU 1 CPU 2
150 =============== ===============
151 { A == 1, B == 2, C = 3, P == &A, Q == &C }
152 B = 4; Q = P;
153 P = &B D = *Q;
155 There is an obvious data dependency here, as the value loaded into D depends on
156 the address retrieved from P by CPU 2. At the end of the sequence, any of the
157 following results are possible:
159 (Q == &A) and (D == 1)
160 (Q == &B) and (D == 2)
161 (Q == &B) and (D == 4)
163 Note that CPU 2 will never try and load C into D because the CPU will load P
164 into Q before issuing the load of *Q.
167 DEVICE OPERATIONS
168 -----------------
170 Some devices present their control interfaces as collections of memory
171 locations, but the order in which the control registers are accessed is very
172 important. For instance, imagine an ethernet card with a set of internal
173 registers that are accessed through an address port register (A) and a data
174 port register (D). To read internal register 5, the following code might then
175 be used:
177 *A = 5;
178 x = *D;
180 but this might show up as either of the following two sequences:
182 STORE *A = 5, x = LOAD *D
183 x = LOAD *D, STORE *A = 5
185 the second of which will almost certainly result in a malfunction, since it set
186 the address _after_ attempting to read the register.
189 GUARANTEES
190 ----------
192 There are some minimal guarantees that may be expected of a CPU:
194 (*) On any given CPU, dependent memory accesses will be issued in order, with
195 respect to itself. This means that for:
197 ACCESS_ONCE(Q) = P; smp_read_barrier_depends(); D = ACCESS_ONCE(*Q);
199 the CPU will issue the following memory operations:
201 Q = LOAD P, D = LOAD *Q
203 and always in that order. On most systems, smp_read_barrier_depends()
204 does nothing, but it is required for DEC Alpha. The ACCESS_ONCE()
205 is required to prevent compiler mischief. Please note that you
206 should normally use something like rcu_dereference() instead of
207 open-coding smp_read_barrier_depends().
209 (*) Overlapping loads and stores within a particular CPU will appear to be
210 ordered within that CPU. This means that for:
212 a = ACCESS_ONCE(*X); ACCESS_ONCE(*X) = b;
214 the CPU will only issue the following sequence of memory operations:
216 a = LOAD *X, STORE *X = b
218 And for:
220 ACCESS_ONCE(*X) = c; d = ACCESS_ONCE(*X);
222 the CPU will only issue:
224 STORE *X = c, d = LOAD *X
226 (Loads and stores overlap if they are targeted at overlapping pieces of
227 memory).
229 And there are a number of things that _must_ or _must_not_ be assumed:
231 (*) It _must_not_ be assumed that the compiler will do what you want with
232 memory references that are not protected by ACCESS_ONCE(). Without
233 ACCESS_ONCE(), the compiler is within its rights to do all sorts
234 of "creative" transformations, which are covered in the Compiler
235 Barrier section.
237 (*) It _must_not_ be assumed that independent loads and stores will be issued
238 in the order given. This means that for:
240 X = *A; Y = *B; *D = Z;
242 we may get any of the following sequences:
244 X = LOAD *A, Y = LOAD *B, STORE *D = Z
245 X = LOAD *A, STORE *D = Z, Y = LOAD *B
246 Y = LOAD *B, X = LOAD *A, STORE *D = Z
247 Y = LOAD *B, STORE *D = Z, X = LOAD *A
248 STORE *D = Z, X = LOAD *A, Y = LOAD *B
249 STORE *D = Z, Y = LOAD *B, X = LOAD *A
251 (*) It _must_ be assumed that overlapping memory accesses may be merged or
252 discarded. This means that for:
254 X = *A; Y = *(A + 4);
256 we may get any one of the following sequences:
258 X = LOAD *A; Y = LOAD *(A + 4);
259 Y = LOAD *(A + 4); X = LOAD *A;
260 {X, Y} = LOAD {*A, *(A + 4) };
262 And for:
264 *A = X; *(A + 4) = Y;
266 we may get any of:
268 STORE *A = X; STORE *(A + 4) = Y;
269 STORE *(A + 4) = Y; STORE *A = X;
270 STORE {*A, *(A + 4) } = {X, Y};
272 And there are anti-guarantees:
274 (*) These guarantees do not apply to bitfields, because compilers often
275 generate code to modify these using non-atomic read-modify-write
276 sequences. Do not attempt to use bitfields to synchronize parallel
277 algorithms.
279 (*) Even in cases where bitfields are protected by locks, all fields
280 in a given bitfield must be protected by one lock. If two fields
281 in a given bitfield are protected by different locks, the compiler's
282 non-atomic read-modify-write sequences can cause an update to one
283 field to corrupt the value of an adjacent field.
285 (*) These guarantees apply only to properly aligned and sized scalar
286 variables. "Properly sized" currently means variables that are
287 the same size as "char", "short", "int" and "long". "Properly
288 aligned" means the natural alignment, thus no constraints for
289 "char", two-byte alignment for "short", four-byte alignment for
290 "int", and either four-byte or eight-byte alignment for "long",
291 on 32-bit and 64-bit systems, respectively. Note that these
292 guarantees were introduced into the C11 standard, so beware when
293 using older pre-C11 compilers (for example, gcc 4.6). The portion
294 of the standard containing this guarantee is Section 3.14, which
295 defines "memory location" as follows:
297 memory location
298 either an object of scalar type, or a maximal sequence
299 of adjacent bit-fields all having nonzero width
301 NOTE 1: Two threads of execution can update and access
302 separate memory locations without interfering with
303 each other.
305 NOTE 2: A bit-field and an adjacent non-bit-field member
306 are in separate memory locations. The same applies
307 to two bit-fields, if one is declared inside a nested
308 structure declaration and the other is not, or if the two
309 are separated by a zero-length bit-field declaration,
310 or if they are separated by a non-bit-field member
311 declaration. It is not safe to concurrently update two
312 bit-fields in the same structure if all members declared
313 between them are also bit-fields, no matter what the
314 sizes of those intervening bit-fields happen to be.
317 =========================
318 WHAT ARE MEMORY BARRIERS?
319 =========================
321 As can be seen above, independent memory operations are effectively performed
322 in random order, but this can be a problem for CPU-CPU interaction and for I/O.
323 What is required is some way of intervening to instruct the compiler and the
324 CPU to restrict the order.
326 Memory barriers are such interventions. They impose a perceived partial
327 ordering over the memory operations on either side of the barrier.
329 Such enforcement is important because the CPUs and other devices in a system
330 can use a variety of tricks to improve performance, including reordering,
331 deferral and combination of memory operations; speculative loads; speculative
332 branch prediction and various types of caching. Memory barriers are used to
333 override or suppress these tricks, allowing the code to sanely control the
334 interaction of multiple CPUs and/or devices.
337 VARIETIES OF MEMORY BARRIER
338 ---------------------------
340 Memory barriers come in four basic varieties:
342 (1) Write (or store) memory barriers.
344 A write memory barrier gives a guarantee that all the STORE operations
345 specified before the barrier will appear to happen before all the STORE
346 operations specified after the barrier with respect to the other
347 components of the system.
349 A write barrier is a partial ordering on stores only; it is not required
350 to have any effect on loads.
352 A CPU can be viewed as committing a sequence of store operations to the
353 memory system as time progresses. All stores before a write barrier will
354 occur in the sequence _before_ all the stores after the write barrier.
356 [!] Note that write barriers should normally be paired with read or data
357 dependency barriers; see the "SMP barrier pairing" subsection.
360 (2) Data dependency barriers.
362 A data dependency barrier is a weaker form of read barrier. In the case
363 where two loads are performed such that the second depends on the result
364 of the first (eg: the first load retrieves the address to which the second
365 load will be directed), a data dependency barrier would be required to
366 make sure that the target of the second load is updated before the address
367 obtained by the first load is accessed.
369 A data dependency barrier is a partial ordering on interdependent loads
370 only; it is not required to have any effect on stores, independent loads
371 or overlapping loads.
373 As mentioned in (1), the other CPUs in the system can be viewed as
374 committing sequences of stores to the memory system that the CPU being
375 considered can then perceive. A data dependency barrier issued by the CPU
376 under consideration guarantees that for any load preceding it, if that
377 load touches one of a sequence of stores from another CPU, then by the
378 time the barrier completes, the effects of all the stores prior to that
379 touched by the load will be perceptible to any loads issued after the data
380 dependency barrier.
382 See the "Examples of memory barrier sequences" subsection for diagrams
383 showing the ordering constraints.
385 [!] Note that the first load really has to have a _data_ dependency and
386 not a control dependency. If the address for the second load is dependent
387 on the first load, but the dependency is through a conditional rather than
388 actually loading the address itself, then it's a _control_ dependency and
389 a full read barrier or better is required. See the "Control dependencies"
390 subsection for more information.
392 [!] Note that data dependency barriers should normally be paired with
393 write barriers; see the "SMP barrier pairing" subsection.
396 (3) Read (or load) memory barriers.
398 A read barrier is a data dependency barrier plus a guarantee that all the
399 LOAD operations specified before the barrier will appear to happen before
400 all the LOAD operations specified after the barrier with respect to the
401 other components of the system.
403 A read barrier is a partial ordering on loads only; it is not required to
404 have any effect on stores.
406 Read memory barriers imply data dependency barriers, and so can substitute
407 for them.
409 [!] Note that read barriers should normally be paired with write barriers;
410 see the "SMP barrier pairing" subsection.
413 (4) General memory barriers.
415 A general memory barrier gives a guarantee that all the LOAD and STORE
416 operations specified before the barrier will appear to happen before all
417 the LOAD and STORE operations specified after the barrier with respect to
418 the other components of the system.
420 A general memory barrier is a partial ordering over both loads and stores.
422 General memory barriers imply both read and write memory barriers, and so
423 can substitute for either.
426 And a couple of implicit varieties:
428 (5) ACQUIRE operations.
430 This acts as a one-way permeable barrier. It guarantees that all memory
431 operations after the ACQUIRE operation will appear to happen after the
432 ACQUIRE operation with respect to the other components of the system.
433 ACQUIRE operations include LOCK operations and smp_load_acquire()
434 operations.
436 Memory operations that occur before an ACQUIRE operation may appear to
437 happen after it completes.
439 An ACQUIRE operation should almost always be paired with a RELEASE
440 operation.
443 (6) RELEASE operations.
445 This also acts as a one-way permeable barrier. It guarantees that all
446 memory operations before the RELEASE operation will appear to happen
447 before the RELEASE operation with respect to the other components of the
448 system. RELEASE operations include UNLOCK operations and
449 smp_store_release() operations.
451 Memory operations that occur after a RELEASE operation may appear to
452 happen before it completes.
454 The use of ACQUIRE and RELEASE operations generally precludes the need
455 for other sorts of memory barrier (but note the exceptions mentioned in
456 the subsection "MMIO write barrier"). In addition, a RELEASE+ACQUIRE
457 pair is -not- guaranteed to act as a full memory barrier. However, after
458 an ACQUIRE on a given variable, all memory accesses preceding any prior
459 RELEASE on that same variable are guaranteed to be visible. In other
460 words, within a given variable's critical section, all accesses of all
461 previous critical sections for that variable are guaranteed to have
462 completed.
464 This means that ACQUIRE acts as a minimal "acquire" operation and
465 RELEASE acts as a minimal "release" operation.
468 Memory barriers are only required where there's a possibility of interaction
469 between two CPUs or between a CPU and a device. If it can be guaranteed that
470 there won't be any such interaction in any particular piece of code, then
471 memory barriers are unnecessary in that piece of code.
474 Note that these are the _minimum_ guarantees. Different architectures may give
475 more substantial guarantees, but they may _not_ be relied upon outside of arch
476 specific code.
479 WHAT MAY NOT BE ASSUMED ABOUT MEMORY BARRIERS?
480 ----------------------------------------------
482 There are certain things that the Linux kernel memory barriers do not guarantee:
484 (*) There is no guarantee that any of the memory accesses specified before a
485 memory barrier will be _complete_ by the completion of a memory barrier
486 instruction; the barrier can be considered to draw a line in that CPU's
487 access queue that accesses of the appropriate type may not cross.
489 (*) There is no guarantee that issuing a memory barrier on one CPU will have
490 any direct effect on another CPU or any other hardware in the system. The
491 indirect effect will be the order in which the second CPU sees the effects
492 of the first CPU's accesses occur, but see the next point:
494 (*) There is no guarantee that a CPU will see the correct order of effects
495 from a second CPU's accesses, even _if_ the second CPU uses a memory
496 barrier, unless the first CPU _also_ uses a matching memory barrier (see
497 the subsection on "SMP Barrier Pairing").
499 (*) There is no guarantee that some intervening piece of off-the-CPU
500 hardware[*] will not reorder the memory accesses. CPU cache coherency
501 mechanisms should propagate the indirect effects of a memory barrier
502 between CPUs, but might not do so in order.
504 [*] For information on bus mastering DMA and coherency please read:
506 Documentation/PCI/pci.txt
507 Documentation/DMA-API-HOWTO.txt
508 Documentation/DMA-API.txt
511 DATA DEPENDENCY BARRIERS
512 ------------------------
514 The usage requirements of data dependency barriers are a little subtle, and
515 it's not always obvious that they're needed. To illustrate, consider the
516 following sequence of events:
518 CPU 1 CPU 2
519 =============== ===============
520 { A == 1, B == 2, C = 3, P == &A, Q == &C }
521 B = 4;
522 <write barrier>
523 ACCESS_ONCE(P) = &B
524 Q = ACCESS_ONCE(P);
525 D = *Q;
527 There's a clear data dependency here, and it would seem that by the end of the
528 sequence, Q must be either &A or &B, and that:
530 (Q == &A) implies (D == 1)
531 (Q == &B) implies (D == 4)
533 But! CPU 2's perception of P may be updated _before_ its perception of B, thus
534 leading to the following situation:
536 (Q == &B) and (D == 2) ????
538 Whilst this may seem like a failure of coherency or causality maintenance, it
539 isn't, and this behaviour can be observed on certain real CPUs (such as the DEC
540 Alpha).
542 To deal with this, a data dependency barrier or better must be inserted
543 between the address load and the data load:
545 CPU 1 CPU 2
546 =============== ===============
547 { A == 1, B == 2, C = 3, P == &A, Q == &C }
548 B = 4;
549 <write barrier>
550 ACCESS_ONCE(P) = &B
551 Q = ACCESS_ONCE(P);
552 <data dependency barrier>
553 D = *Q;
555 This enforces the occurrence of one of the two implications, and prevents the
556 third possibility from arising.
558 [!] Note that this extremely counterintuitive situation arises most easily on
559 machines with split caches, so that, for example, one cache bank processes
560 even-numbered cache lines and the other bank processes odd-numbered cache
561 lines. The pointer P might be stored in an odd-numbered cache line, and the
562 variable B might be stored in an even-numbered cache line. Then, if the
563 even-numbered bank of the reading CPU's cache is extremely busy while the
564 odd-numbered bank is idle, one can see the new value of the pointer P (&B),
565 but the old value of the variable B (2).
568 Another example of where data dependency barriers might be required is where a
569 number is read from memory and then used to calculate the index for an array
570 access:
572 CPU 1 CPU 2
573 =============== ===============
574 { M[0] == 1, M[1] == 2, M[3] = 3, P == 0, Q == 3 }
575 M[1] = 4;
576 <write barrier>
577 ACCESS_ONCE(P) = 1
578 Q = ACCESS_ONCE(P);
579 <data dependency barrier>
580 D = M[Q];
583 The data dependency barrier is very important to the RCU system,
584 for example. See rcu_assign_pointer() and rcu_dereference() in
585 include/linux/rcupdate.h. This permits the current target of an RCU'd
586 pointer to be replaced with a new modified target, without the replacement
587 target appearing to be incompletely initialised.
589 See also the subsection on "Cache Coherency" for a more thorough example.
592 CONTROL DEPENDENCIES
593 --------------------
595 A control dependency requires a full read memory barrier, not simply a data
596 dependency barrier to make it work correctly. Consider the following bit of
597 code:
599 q = ACCESS_ONCE(a);
600 if (q) {
601 <data dependency barrier> /* BUG: No data dependency!!! */
602 p = ACCESS_ONCE(b);
603 }
605 This will not have the desired effect because there is no actual data
606 dependency, but rather a control dependency that the CPU may short-circuit
607 by attempting to predict the outcome in advance, so that other CPUs see
608 the load from b as having happened before the load from a. In such a
609 case what's actually required is:
611 q = ACCESS_ONCE(a);
612 if (q) {
613 <read barrier>
614 p = ACCESS_ONCE(b);
615 }
617 However, stores are not speculated. This means that ordering -is- provided
618 in the following example:
620 q = ACCESS_ONCE(a);
621 if (q) {
622 ACCESS_ONCE(b) = p;
623 }
625 Please note that ACCESS_ONCE() is not optional! Without the
626 ACCESS_ONCE(), might combine the load from 'a' with other loads from
627 'a', and the store to 'b' with other stores to 'b', with possible highly
628 counterintuitive effects on ordering.
630 Worse yet, if the compiler is able to prove (say) that the value of
631 variable 'a' is always non-zero, it would be well within its rights
632 to optimize the original example by eliminating the "if" statement
633 as follows:
635 q = a;
636 b = p; /* BUG: Compiler and CPU can both reorder!!! */
638 So don't leave out the ACCESS_ONCE().
640 It is tempting to try to enforce ordering on identical stores on both
641 branches of the "if" statement as follows:
643 q = ACCESS_ONCE(a);
644 if (q) {
645 barrier();
646 ACCESS_ONCE(b) = p;
647 do_something();
648 } else {
649 barrier();
650 ACCESS_ONCE(b) = p;
651 do_something_else();
652 }
654 Unfortunately, current compilers will transform this as follows at high
655 optimization levels:
657 q = ACCESS_ONCE(a);
658 barrier();
659 ACCESS_ONCE(b) = p; /* BUG: No ordering vs. load from a!!! */
660 if (q) {
661 /* ACCESS_ONCE(b) = p; -- moved up, BUG!!! */
662 do_something();
663 } else {
664 /* ACCESS_ONCE(b) = p; -- moved up, BUG!!! */
665 do_something_else();
666 }
668 Now there is no conditional between the load from 'a' and the store to
669 'b', which means that the CPU is within its rights to reorder them:
670 The conditional is absolutely required, and must be present in the
671 assembly code even after all compiler optimizations have been applied.
672 Therefore, if you need ordering in this example, you need explicit
673 memory barriers, for example, smp_store_release():
675 q = ACCESS_ONCE(a);
676 if (q) {
677 smp_store_release(&b, p);
678 do_something();
679 } else {
680 smp_store_release(&b, p);
681 do_something_else();
682 }
684 In contrast, without explicit memory barriers, two-legged-if control
685 ordering is guaranteed only when the stores differ, for example:
687 q = ACCESS_ONCE(a);
688 if (q) {
689 ACCESS_ONCE(b) = p;
690 do_something();
691 } else {
692 ACCESS_ONCE(b) = r;
693 do_something_else();
694 }
696 The initial ACCESS_ONCE() is still required to prevent the compiler from
697 proving the value of 'a'.
699 In addition, you need to be careful what you do with the local variable 'q',
700 otherwise the compiler might be able to guess the value and again remove
701 the needed conditional. For example:
703 q = ACCESS_ONCE(a);
704 if (q % MAX) {
705 ACCESS_ONCE(b) = p;
706 do_something();
707 } else {
708 ACCESS_ONCE(b) = r;
709 do_something_else();
710 }
712 If MAX is defined to be 1, then the compiler knows that (q % MAX) is
713 equal to zero, in which case the compiler is within its rights to
714 transform the above code into the following:
716 q = ACCESS_ONCE(a);
717 ACCESS_ONCE(b) = p;
718 do_something_else();
720 Given this transformation, the CPU is not required to respect the ordering
721 between the load from variable 'a' and the store to variable 'b'. It is
722 tempting to add a barrier(), but this does not help. The conditional
723 is gone, and the barrier won't bring it back. Therefore, if you are
724 relying on this ordering, you should make sure that MAX is greater than
725 one, perhaps as follows:
727 q = ACCESS_ONCE(a);
728 BUILD_BUG_ON(MAX <= 1); /* Order load from a with store to b. */
729 if (q % MAX) {
730 ACCESS_ONCE(b) = p;
731 do_something();
732 } else {
733 ACCESS_ONCE(b) = r;
734 do_something_else();
735 }
737 Please note once again that the stores to 'b' differ. If they were
738 identical, as noted earlier, the compiler could pull this store outside
739 of the 'if' statement.
741 You must also be careful not to rely too much on boolean short-circuit
742 evaluation. Consider this example:
744 q = ACCESS_ONCE(a);
745 if (a || 1 > 0)
746 ACCESS_ONCE(b) = 1;
748 Because the second condition is always true, the compiler can transform
749 this example as following, defeating control dependency:
751 q = ACCESS_ONCE(a);
752 ACCESS_ONCE(b) = 1;
754 This example underscores the need to ensure that the compiler cannot
755 out-guess your code. More generally, although ACCESS_ONCE() does force
756 the compiler to actually emit code for a given load, it does not force
757 the compiler to use the results.
759 Finally, control dependencies do -not- provide transitivity. This is
760 demonstrated by two related examples, with the initial values of
761 x and y both being zero:
763 CPU 0 CPU 1
764 ===================== =====================
765 r1 = ACCESS_ONCE(x); r2 = ACCESS_ONCE(y);
766 if (r1 > 0) if (r2 > 0)
767 ACCESS_ONCE(y) = 1; ACCESS_ONCE(x) = 1;
769 assert(!(r1 == 1 && r2 == 1));
771 The above two-CPU example will never trigger the assert(). However,
772 if control dependencies guaranteed transitivity (which they do not),
773 then adding the following CPU would guarantee a related assertion:
775 CPU 2
776 =====================
777 ACCESS_ONCE(x) = 2;
779 assert(!(r1 == 2 && r2 == 1 && x == 2)); /* FAILS!!! */
781 But because control dependencies do -not- provide transitivity, the above
782 assertion can fail after the combined three-CPU example completes. If you
783 need the three-CPU example to provide ordering, you will need smp_mb()
784 between the loads and stores in the CPU 0 and CPU 1 code fragments,
785 that is, just before or just after the "if" statements.
787 These two examples are the LB and WWC litmus tests from this paper:
788 http://www.cl.cam.ac.uk/users/pes20/ppc-supplemental/test6.pdf and this
789 site: https://www.cl.cam.ac.uk/~pes20/ppcmem/index.html.
791 In summary:
793 (*) Control dependencies can order prior loads against later stores.
794 However, they do -not- guarantee any other sort of ordering:
795 Not prior loads against later loads, nor prior stores against
796 later anything. If you need these other forms of ordering,
797 use smp_rmb(), smp_wmb(), or, in the case of prior stores and
798 later loads, smp_mb().
800 (*) If both legs of the "if" statement begin with identical stores
801 to the same variable, a barrier() statement is required at the
802 beginning of each leg of the "if" statement.
804 (*) Control dependencies require at least one run-time conditional
805 between the prior load and the subsequent store, and this
806 conditional must involve the prior load. If the compiler
807 is able to optimize the conditional away, it will have also
808 optimized away the ordering. Careful use of ACCESS_ONCE() can
809 help to preserve the needed conditional.
811 (*) Control dependencies require that the compiler avoid reordering the
812 dependency into nonexistence. Careful use of ACCESS_ONCE() or
813 barrier() can help to preserve your control dependency. Please
814 see the Compiler Barrier section for more information.
816 (*) Control dependencies do -not- provide transitivity. If you
817 need transitivity, use smp_mb().
820 SMP BARRIER PAIRING
821 -------------------
823 When dealing with CPU-CPU interactions, certain types of memory barrier should
824 always be paired. A lack of appropriate pairing is almost certainly an error.
826 General barriers pair with each other, though they also pair with
827 most other types of barriers, albeit without transitivity. An acquire
828 barrier pairs with a release barrier, but both may also pair with other
829 barriers, including of course general barriers. A write barrier pairs
830 with a data dependency barrier, an acquire barrier, a release barrier,
831 a read barrier, or a general barrier. Similarly a read barrier or a
832 data dependency barrier pairs with a write barrier, an acquire barrier,
833 a release barrier, or a general barrier:
835 CPU 1 CPU 2
836 =============== ===============
837 ACCESS_ONCE(a) = 1;
838 <write barrier>
839 ACCESS_ONCE(b) = 2; x = ACCESS_ONCE(b);
840 <read barrier>
841 y = ACCESS_ONCE(a);
843 Or:
845 CPU 1 CPU 2
846 =============== ===============================
847 a = 1;
848 <write barrier>
849 ACCESS_ONCE(b) = &a; x = ACCESS_ONCE(b);
850 <data dependency barrier>
851 y = *x;
853 Basically, the read barrier always has to be there, even though it can be of
854 the "weaker" type.
856 [!] Note that the stores before the write barrier would normally be expected to
857 match the loads after the read barrier or the data dependency barrier, and vice
858 versa:
860 CPU 1 CPU 2
861 =================== ===================
862 ACCESS_ONCE(a) = 1; }---- --->{ v = ACCESS_ONCE(c);
863 ACCESS_ONCE(b) = 2; } \ / { w = ACCESS_ONCE(d);
864 <write barrier> \ <read barrier>
865 ACCESS_ONCE(c) = 3; } / \ { x = ACCESS_ONCE(a);
866 ACCESS_ONCE(d) = 4; }---- --->{ y = ACCESS_ONCE(b);
869 EXAMPLES OF MEMORY BARRIER SEQUENCES
870 ------------------------------------
872 Firstly, write barriers act as partial orderings on store operations.
873 Consider the following sequence of events:
875 CPU 1
876 =======================
877 STORE A = 1
878 STORE B = 2
879 STORE C = 3
880 <write barrier>
881 STORE D = 4
882 STORE E = 5
884 This sequence of events is committed to the memory coherence system in an order
885 that the rest of the system might perceive as the unordered set of { STORE A,
886 STORE B, STORE C } all occurring before the unordered set of { STORE D, STORE E
887 }:
889 +-------+ : :
890 | | +------+
891 | |------>| C=3 | } /\
892 | | : +------+ }----- \ -----> Events perceptible to
893 | | : | A=1 | } \/ the rest of the system
894 | | : +------+ }
895 | CPU 1 | : | B=2 | }
896 | | +------+ }
897 | | wwwwwwwwwwwwwwww } <--- At this point the write barrier
898 | | +------+ } requires all stores prior to the
899 | | : | E=5 | } barrier to be committed before
900 | | : +------+ } further stores may take place
901 | |------>| D=4 | }
902 | | +------+
903 +-------+ : :
904 |
905 | Sequence in which stores are committed to the
906 | memory system by CPU 1
907 V
910 Secondly, data dependency barriers act as partial orderings on data-dependent
911 loads. Consider the following sequence of events:
913 CPU 1 CPU 2
914 ======================= =======================
915 { B = 7; X = 9; Y = 8; C = &Y }
916 STORE A = 1
917 STORE B = 2
918 <write barrier>
919 STORE C = &B LOAD X
920 STORE D = 4 LOAD C (gets &B)
921 LOAD *C (reads B)
923 Without intervention, CPU 2 may perceive the events on CPU 1 in some
924 effectively random order, despite the write barrier issued by CPU 1:
926 +-------+ : : : :
927 | | +------+ +-------+ | Sequence of update
928 | |------>| B=2 |----- --->| Y->8 | | of perception on
929 | | : +------+ \ +-------+ | CPU 2
930 | CPU 1 | : | A=1 | \ --->| C->&Y | V
931 | | +------+ | +-------+
932 | | wwwwwwwwwwwwwwww | : :
933 | | +------+ | : :
934 | | : | C=&B |--- | : : +-------+
935 | | : +------+ \ | +-------+ | |
936 | |------>| D=4 | ----------->| C->&B |------>| |
937 | | +------+ | +-------+ | |
938 +-------+ : : | : : | |
939 | : : | |
940 | : : | CPU 2 |
941 | +-------+ | |
942 Apparently incorrect ---> | | B->7 |------>| |
943 perception of B (!) | +-------+ | |
944 | : : | |
945 | +-------+ | |
946 The load of X holds ---> \ | X->9 |------>| |
947 up the maintenance \ +-------+ | |
948 of coherence of B ----->| B->2 | +-------+
949 +-------+
950 : :
953 In the above example, CPU 2 perceives that B is 7, despite the load of *C
954 (which would be B) coming after the LOAD of C.
956 If, however, a data dependency barrier were to be placed between the load of C
957 and the load of *C (ie: B) on CPU 2:
959 CPU 1 CPU 2
960 ======================= =======================
961 { B = 7; X = 9; Y = 8; C = &Y }
962 STORE A = 1
963 STORE B = 2
964 <write barrier>
965 STORE C = &B LOAD X
966 STORE D = 4 LOAD C (gets &B)
967 <data dependency barrier>
968 LOAD *C (reads B)
970 then the following will occur:
972 +-------+ : : : :
973 | | +------+ +-------+
974 | |------>| B=2 |----- --->| Y->8 |
975 | | : +------+ \ +-------+
976 | CPU 1 | : | A=1 | \ --->| C->&Y |
977 | | +------+ | +-------+
978 | | wwwwwwwwwwwwwwww | : :
979 | | +------+ | : :
980 | | : | C=&B |--- | : : +-------+
981 | | : +------+ \ | +-------+ | |
982 | |------>| D=4 | ----------->| C->&B |------>| |
983 | | +------+ | +-------+ | |
984 +-------+ : : | : : | |
985 | : : | |
986 | : : | CPU 2 |
987 | +-------+ | |
988 | | X->9 |------>| |
989 | +-------+ | |
990 Makes sure all effects ---> \ ddddddddddddddddd | |
991 prior to the store of C \ +-------+ | |
992 are perceptible to ----->| B->2 |------>| |
993 subsequent loads +-------+ | |
994 : : +-------+
997 And thirdly, a read barrier acts as a partial order on loads. Consider the
998 following sequence of events:
1000 CPU 1 CPU 2
1001 ======================= =======================
1002 { A = 0, B = 9 }
1003 STORE A=1
1004 <write barrier>
1005 STORE B=2
1006 LOAD B
1007 LOAD A
1009 Without intervention, CPU 2 may then choose to perceive the events on CPU 1 in
1010 some effectively random order, despite the write barrier issued by CPU 1:
1012 +-------+ : : : :
1013 | | +------+ +-------+
1014 | |------>| A=1 |------ --->| A->0 |
1015 | | +------+ \ +-------+
1016 | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
1017 | | +------+ | +-------+
1018 | |------>| B=2 |--- | : :
1019 | | +------+ \ | : : +-------+
1020 +-------+ : : \ | +-------+ | |
1021 ---------->| B->2 |------>| |
1022 | +-------+ | CPU 2 |
1023 | | A->0 |------>| |
1024 | +-------+ | |
1025 | : : +-------+
1026 \ : :
1027 \ +-------+
1028 ---->| A->1 |
1029 +-------+
1030 : :
1033 If, however, a read barrier were to be placed between the load of B and the
1034 load of A on CPU 2:
1036 CPU 1 CPU 2
1037 ======================= =======================
1038 { A = 0, B = 9 }
1039 STORE A=1
1040 <write barrier>
1041 STORE B=2
1042 LOAD B
1043 <read barrier>
1044 LOAD A
1046 then the partial ordering imposed by CPU 1 will be perceived correctly by CPU
1047 2:
1049 +-------+ : : : :
1050 | | +------+ +-------+
1051 | |------>| A=1 |------ --->| A->0 |
1052 | | +------+ \ +-------+
1053 | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
1054 | | +------+ | +-------+
1055 | |------>| B=2 |--- | : :
1056 | | +------+ \ | : : +-------+
1057 +-------+ : : \ | +-------+ | |
1058 ---------->| B->2 |------>| |
1059 | +-------+ | CPU 2 |
1060 | : : | |
1061 | : : | |
1062 At this point the read ----> \ rrrrrrrrrrrrrrrrr | |
1063 barrier causes all effects \ +-------+ | |
1064 prior to the storage of B ---->| A->1 |------>| |
1065 to be perceptible to CPU 2 +-------+ | |
1066 : : +-------+
1069 To illustrate this more completely, consider what could happen if the code
1070 contained a load of A either side of the read barrier:
1072 CPU 1 CPU 2
1073 ======================= =======================
1074 { A = 0, B = 9 }
1075 STORE A=1
1076 <write barrier>
1077 STORE B=2
1078 LOAD B
1079 LOAD A [first load of A]
1080 <read barrier>
1081 LOAD A [second load of A]
1083 Even though the two loads of A both occur after the load of B, they may both
1084 come up with different values:
1086 +-------+ : : : :
1087 | | +------+ +-------+
1088 | |------>| A=1 |------ --->| A->0 |
1089 | | +------+ \ +-------+
1090 | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
1091 | | +------+ | +-------+
1092 | |------>| B=2 |--- | : :
1093 | | +------+ \ | : : +-------+
1094 +-------+ : : \ | +-------+ | |
1095 ---------->| B->2 |------>| |
1096 | +-------+ | CPU 2 |
1097 | : : | |
1098 | : : | |
1099 | +-------+ | |
1100 | | A->0 |------>| 1st |
1101 | +-------+ | |
1102 At this point the read ----> \ rrrrrrrrrrrrrrrrr | |
1103 barrier causes all effects \ +-------+ | |
1104 prior to the storage of B ---->| A->1 |------>| 2nd |
1105 to be perceptible to CPU 2 +-------+ | |
1106 : : +-------+
1109 But it may be that the update to A from CPU 1 becomes perceptible to CPU 2
1110 before the read barrier completes anyway:
1112 +-------+ : : : :
1113 | | +------+ +-------+
1114 | |------>| A=1 |------ --->| A->0 |
1115 | | +------+ \ +-------+
1116 | CPU 1 | wwwwwwwwwwwwwwww \ --->| B->9 |
1117 | | +------+ | +-------+
1118 | |------>| B=2 |--- | : :
1119 | | +------+ \ | : : +-------+
1120 +-------+ : : \ | +-------+ | |
1121 ---------->| B->2 |------>| |
1122 | +-------+ | CPU 2 |
1123 | : : | |
1124 \ : : | |
1125 \ +-------+ | |
1126 ---->| A->1 |------>| 1st |
1127 +-------+ | |
1128 rrrrrrrrrrrrrrrrr | |
1129 +-------+ | |
1130 | A->1 |------>| 2nd |
1131 +-------+ | |
1132 : : +-------+
1135 The guarantee is that the second load will always come up with A == 1 if the
1136 load of B came up with B == 2. No such guarantee exists for the first load of
1137 A; that may come up with either A == 0 or A == 1.
1140 READ MEMORY BARRIERS VS LOAD SPECULATION
1141 ----------------------------------------
1143 Many CPUs speculate with loads: that is they see that they will need to load an
1144 item from memory, and they find a time where they're not using the bus for any
1145 other loads, and so do the load in advance - even though they haven't actually
1146 got to that point in the instruction execution flow yet. This permits the
1147 actual load instruction to potentially complete immediately because the CPU
1148 already has the value to hand.
1150 It may turn out that the CPU didn't actually need the value - perhaps because a
1151 branch circumvented the load - in which case it can discard the value or just
1152 cache it for later use.
1154 Consider:
1156 CPU 1 CPU 2
1157 ======================= =======================
1158 LOAD B
1159 DIVIDE } Divide instructions generally
1160 DIVIDE } take a long time to perform
1161 LOAD A
1163 Which might appear as this:
1165 : : +-------+
1166 +-------+ | |
1167 --->| B->2 |------>| |
1168 +-------+ | CPU 2 |
1169 : :DIVIDE | |
1170 +-------+ | |
1171 The CPU being busy doing a ---> --->| A->0 |~~~~ | |
1172 division speculates on the +-------+ ~ | |
1173 LOAD of A : : ~ | |
1174 : :DIVIDE | |
1175 : : ~ | |
1176 Once the divisions are complete --> : : ~-->| |
1177 the CPU can then perform the : : | |
1178 LOAD with immediate effect : : +-------+
1181 Placing a read barrier or a data dependency barrier just before the second
1182 load:
1184 CPU 1 CPU 2
1185 ======================= =======================
1186 LOAD B
1187 DIVIDE
1188 DIVIDE
1189 <read barrier>
1190 LOAD A
1192 will force any value speculatively obtained to be reconsidered to an extent
1193 dependent on the type of barrier used. If there was no change made to the
1194 speculated memory location, then the speculated value will just be used:
1196 : : +-------+
1197 +-------+ | |
1198 --->| B->2 |------>| |
1199 +-------+ | CPU 2 |
1200 : :DIVIDE | |
1201 +-------+ | |
1202 The CPU being busy doing a ---> --->| A->0 |~~~~ | |
1203 division speculates on the +-------+ ~ | |
1204 LOAD of A : : ~ | |
1205 : :DIVIDE | |
1206 : : ~ | |
1207 : : ~ | |
1208 rrrrrrrrrrrrrrrr~ | |
1209 : : ~ | |
1210 : : ~-->| |
1211 : : | |
1212 : : +-------+
1215 but if there was an update or an invalidation from another CPU pending, then
1216 the speculation will be cancelled and the value reloaded:
1218 : : +-------+
1219 +-------+ | |
1220 --->| B->2 |------>| |
1221 +-------+ | CPU 2 |
1222 : :DIVIDE | |
1223 +-------+ | |
1224 The CPU being busy doing a ---> --->| A->0 |~~~~ | |
1225 division speculates on the +-------+ ~ | |
1226 LOAD of A : : ~ | |
1227 : :DIVIDE | |
1228 : : ~ | |
1229 : : ~ | |
1230 rrrrrrrrrrrrrrrrr | |
1231 +-------+ | |
1232 The speculation is discarded ---> --->| A->1 |------>| |
1233 and an updated value is +-------+ | |
1234 retrieved : : +-------+
1237 TRANSITIVITY
1238 ------------
1240 Transitivity is a deeply intuitive notion about ordering that is not
1241 always provided by real computer systems. The following example
1242 demonstrates transitivity (also called "cumulativity"):
1244 CPU 1 CPU 2 CPU 3
1245 ======================= ======================= =======================
1246 { X = 0, Y = 0 }
1247 STORE X=1 LOAD X STORE Y=1
1248 <general barrier> <general barrier>
1249 LOAD Y LOAD X
1251 Suppose that CPU 2's load from X returns 1 and its load from Y returns 0.
1252 This indicates that CPU 2's load from X in some sense follows CPU 1's
1253 store to X and that CPU 2's load from Y in some sense preceded CPU 3's
1254 store to Y. The question is then "Can CPU 3's load from X return 0?"
1256 Because CPU 2's load from X in some sense came after CPU 1's store, it
1257 is natural to expect that CPU 3's load from X must therefore return 1.
1258 This expectation is an example of transitivity: if a load executing on
1259 CPU A follows a load from the same variable executing on CPU B, then
1260 CPU A's load must either return the same value that CPU B's load did,
1261 or must return some later value.
1263 In the Linux kernel, use of general memory barriers guarantees
1264 transitivity. Therefore, in the above example, if CPU 2's load from X
1265 returns 1 and its load from Y returns 0, then CPU 3's load from X must
1266 also return 1.
1268 However, transitivity is -not- guaranteed for read or write barriers.
1269 For example, suppose that CPU 2's general barrier in the above example
1270 is changed to a read barrier as shown below:
1272 CPU 1 CPU 2 CPU 3
1273 ======================= ======================= =======================
1274 { X = 0, Y = 0 }
1275 STORE X=1 LOAD X STORE Y=1
1276 <read barrier> <general barrier>
1277 LOAD Y LOAD X
1279 This substitution destroys transitivity: in this example, it is perfectly
1280 legal for CPU 2's load from X to return 1, its load from Y to return 0,
1281 and CPU 3's load from X to return 0.
1283 The key point is that although CPU 2's read barrier orders its pair
1284 of loads, it does not guarantee to order CPU 1's store. Therefore, if
1285 this example runs on a system where CPUs 1 and 2 share a store buffer
1286 or a level of cache, CPU 2 might have early access to CPU 1's writes.
1287 General barriers are therefore required to ensure that all CPUs agree
1288 on the combined order of CPU 1's and CPU 2's accesses.
1290 To reiterate, if your code requires transitivity, use general barriers
1291 throughout.
1294 ========================
1295 EXPLICIT KERNEL BARRIERS
1296 ========================
1298 The Linux kernel has a variety of different barriers that act at different
1299 levels:
1301 (*) Compiler barrier.
1303 (*) CPU memory barriers.
1305 (*) MMIO write barrier.
1308 COMPILER BARRIER
1309 ----------------
1311 The Linux kernel has an explicit compiler barrier function that prevents the
1312 compiler from moving the memory accesses either side of it to the other side:
1314 barrier();
1316 This is a general barrier -- there are no read-read or write-write variants
1317 of barrier(). However, ACCESS_ONCE() can be thought of as a weak form
1318 for barrier() that affects only the specific accesses flagged by the
1319 ACCESS_ONCE().
1321 The barrier() function has the following effects:
1323 (*) Prevents the compiler from reordering accesses following the
1324 barrier() to precede any accesses preceding the barrier().
1325 One example use for this property is to ease communication between
1326 interrupt-handler code and the code that was interrupted.
1328 (*) Within a loop, forces the compiler to load the variables used
1329 in that loop's conditional on each pass through that loop.
1331 The ACCESS_ONCE() function can prevent any number of optimizations that,
1332 while perfectly safe in single-threaded code, can be fatal in concurrent
1333 code. Here are some examples of these sorts of optimizations:
1335 (*) The compiler is within its rights to reorder loads and stores
1336 to the same variable, and in some cases, the CPU is within its
1337 rights to reorder loads to the same variable. This means that
1338 the following code:
1340 a[0] = x;
1341 a[1] = x;
1343 Might result in an older value of x stored in a[1] than in a[0].
1344 Prevent both the compiler and the CPU from doing this as follows:
1346 a[0] = ACCESS_ONCE(x);
1347 a[1] = ACCESS_ONCE(x);
1349 In short, ACCESS_ONCE() provides cache coherence for accesses from
1350 multiple CPUs to a single variable.
1352 (*) The compiler is within its rights to merge successive loads from
1353 the same variable. Such merging can cause the compiler to "optimize"
1354 the following code:
1356 while (tmp = a)
1357 do_something_with(tmp);
1359 into the following code, which, although in some sense legitimate
1360 for single-threaded code, is almost certainly not what the developer
1361 intended:
1363 if (tmp = a)
1364 for (;;)
1365 do_something_with(tmp);
1367 Use ACCESS_ONCE() to prevent the compiler from doing this to you:
1369 while (tmp = ACCESS_ONCE(a))
1370 do_something_with(tmp);
1372 (*) The compiler is within its rights to reload a variable, for example,
1373 in cases where high register pressure prevents the compiler from
1374 keeping all data of interest in registers. The compiler might
1375 therefore optimize the variable 'tmp' out of our previous example:
1377 while (tmp = a)
1378 do_something_with(tmp);
1380 This could result in the following code, which is perfectly safe in
1381 single-threaded code, but can be fatal in concurrent code:
1383 while (a)
1384 do_something_with(a);
1386 For example, the optimized version of this code could result in
1387 passing a zero to do_something_with() in the case where the variable
1388 a was modified by some other CPU between the "while" statement and
1389 the call to do_something_with().
1391 Again, use ACCESS_ONCE() to prevent the compiler from doing this:
1393 while (tmp = ACCESS_ONCE(a))
1394 do_something_with(tmp);
1396 Note that if the compiler runs short of registers, it might save
1397 tmp onto the stack. The overhead of this saving and later restoring
1398 is why compilers reload variables. Doing so is perfectly safe for
1399 single-threaded code, so you need to tell the compiler about cases
1400 where it is not safe.
1402 (*) The compiler is within its rights to omit a load entirely if it knows
1403 what the value will be. For example, if the compiler can prove that
1404 the value of variable 'a' is always zero, it can optimize this code:
1406 while (tmp = a)
1407 do_something_with(tmp);
1409 Into this:
1411 do { } while (0);
1413 This transformation is a win for single-threaded code because it gets
1414 rid of a load and a branch. The problem is that the compiler will
1415 carry out its proof assuming that the current CPU is the only one
1416 updating variable 'a'. If variable 'a' is shared, then the compiler's
1417 proof will be erroneous. Use ACCESS_ONCE() to tell the compiler
1418 that it doesn't know as much as it thinks it does:
1420 while (tmp = ACCESS_ONCE(a))
1421 do_something_with(tmp);
1423 But please note that the compiler is also closely watching what you
1424 do with the value after the ACCESS_ONCE(). For example, suppose you
1425 do the following and MAX is a preprocessor macro with the value 1:
1427 while ((tmp = ACCESS_ONCE(a)) % MAX)
1428 do_something_with(tmp);
1430 Then the compiler knows that the result of the "%" operator applied
1431 to MAX will always be zero, again allowing the compiler to optimize
1432 the code into near-nonexistence. (It will still load from the
1433 variable 'a'.)
1435 (*) Similarly, the compiler is within its rights to omit a store entirely
1436 if it knows that the variable already has the value being stored.
1437 Again, the compiler assumes that the current CPU is the only one
1438 storing into the variable, which can cause the compiler to do the
1439 wrong thing for shared variables. For example, suppose you have
1440 the following:
1442 a = 0;
1443 /* Code that does not store to variable a. */
1444 a = 0;
1446 The compiler sees that the value of variable 'a' is already zero, so
1447 it might well omit the second store. This would come as a fatal
1448 surprise if some other CPU might have stored to variable 'a' in the
1449 meantime.
1451 Use ACCESS_ONCE() to prevent the compiler from making this sort of
1452 wrong guess:
1454 ACCESS_ONCE(a) = 0;
1455 /* Code that does not store to variable a. */
1456 ACCESS_ONCE(a) = 0;
1458 (*) The compiler is within its rights to reorder memory accesses unless
1459 you tell it not to. For example, consider the following interaction
1460 between process-level code and an interrupt handler:
1462 void process_level(void)
1463 {
1464 msg = get_message();
1465 flag = true;
1466 }
1468 void interrupt_handler(void)
1469 {
1470 if (flag)
1471 process_message(msg);
1472 }
1474 There is nothing to prevent the compiler from transforming
1475 process_level() to the following, in fact, this might well be a
1476 win for single-threaded code:
1478 void process_level(void)
1479 {
1480 flag = true;
1481 msg = get_message();
1482 }
1484 If the interrupt occurs between these two statement, then
1485 interrupt_handler() might be passed a garbled msg. Use ACCESS_ONCE()
1486 to prevent this as follows:
1488 void process_level(void)
1489 {
1490 ACCESS_ONCE(msg) = get_message();
1491 ACCESS_ONCE(flag) = true;
1492 }
1494 void interrupt_handler(void)
1495 {
1496 if (ACCESS_ONCE(flag))
1497 process_message(ACCESS_ONCE(msg));
1498 }
1500 Note that the ACCESS_ONCE() wrappers in interrupt_handler()
1501 are needed if this interrupt handler can itself be interrupted
1502 by something that also accesses 'flag' and 'msg', for example,
1503 a nested interrupt or an NMI. Otherwise, ACCESS_ONCE() is not
1504 needed in interrupt_handler() other than for documentation purposes.
1505 (Note also that nested interrupts do not typically occur in modern
1506 Linux kernels, in fact, if an interrupt handler returns with
1507 interrupts enabled, you will get a WARN_ONCE() splat.)
1509 You should assume that the compiler can move ACCESS_ONCE() past
1510 code not containing ACCESS_ONCE(), barrier(), or similar primitives.
1512 This effect could also be achieved using barrier(), but ACCESS_ONCE()
1513 is more selective: With ACCESS_ONCE(), the compiler need only forget
1514 the contents of the indicated memory locations, while with barrier()
1515 the compiler must discard the value of all memory locations that
1516 it has currented cached in any machine registers. Of course,
1517 the compiler must also respect the order in which the ACCESS_ONCE()s
1518 occur, though the CPU of course need not do so.
1520 (*) The compiler is within its rights to invent stores to a variable,
1521 as in the following example:
1523 if (a)
1524 b = a;
1525 else
1526 b = 42;
1528 The compiler might save a branch by optimizing this as follows:
1530 b = 42;
1531 if (a)
1532 b = a;
1534 In single-threaded code, this is not only safe, but also saves
1535 a branch. Unfortunately, in concurrent code, this optimization
1536 could cause some other CPU to see a spurious value of 42 -- even
1537 if variable 'a' was never zero -- when loading variable 'b'.
1538 Use ACCESS_ONCE() to prevent this as follows:
1540 if (a)
1541 ACCESS_ONCE(b) = a;
1542 else
1543 ACCESS_ONCE(b) = 42;
1545 The compiler can also invent loads. These are usually less
1546 damaging, but they can result in cache-line bouncing and thus in
1547 poor performance and scalability. Use ACCESS_ONCE() to prevent
1548 invented loads.
1550 (*) For aligned memory locations whose size allows them to be accessed
1551 with a single memory-reference instruction, prevents "load tearing"
1552 and "store tearing," in which a single large access is replaced by
1553 multiple smaller accesses. For example, given an architecture having
1554 16-bit store instructions with 7-bit immediate fields, the compiler
1555 might be tempted to use two 16-bit store-immediate instructions to
1556 implement the following 32-bit store:
1558 p = 0x00010002;
1560 Please note that GCC really does use this sort of optimization,
1561 which is not surprising given that it would likely take more
1562 than two instructions to build the constant and then store it.
1563 This optimization can therefore be a win in single-threaded code.
1564 In fact, a recent bug (since fixed) caused GCC to incorrectly use
1565 this optimization in a volatile store. In the absence of such bugs,
1566 use of ACCESS_ONCE() prevents store tearing in the following example:
1568 ACCESS_ONCE(p) = 0x00010002;
1570 Use of packed structures can also result in load and store tearing,
1571 as in this example:
1573 struct __attribute__((__packed__)) foo {
1574 short a;
1575 int b;
1576 short c;
1577 };
1578 struct foo foo1, foo2;
1579 ...
1581 foo2.a = foo1.a;
1582 foo2.b = foo1.b;
1583 foo2.c = foo1.c;
1585 Because there are no ACCESS_ONCE() wrappers and no volatile markings,
1586 the compiler would be well within its rights to implement these three
1587 assignment statements as a pair of 32-bit loads followed by a pair
1588 of 32-bit stores. This would result in load tearing on 'foo1.b'
1589 and store tearing on 'foo2.b'. ACCESS_ONCE() again prevents tearing
1590 in this example:
1592 foo2.a = foo1.a;
1593 ACCESS_ONCE(foo2.b) = ACCESS_ONCE(foo1.b);
1594 foo2.c = foo1.c;
1596 All that aside, it is never necessary to use ACCESS_ONCE() on a variable
1597 that has been marked volatile. For example, because 'jiffies' is marked
1598 volatile, it is never necessary to say ACCESS_ONCE(jiffies). The reason
1599 for this is that ACCESS_ONCE() is implemented as a volatile cast, which
1600 has no effect when its argument is already marked volatile.
1602 Please note that these compiler barriers have no direct effect on the CPU,
1603 which may then reorder things however it wishes.
1606 CPU MEMORY BARRIERS
1607 -------------------
1609 The Linux kernel has eight basic CPU memory barriers:
1611 TYPE MANDATORY SMP CONDITIONAL
1612 =============== ======================= ===========================
1613 GENERAL mb() smp_mb()
1614 WRITE wmb() smp_wmb()
1615 READ rmb() smp_rmb()
1616 DATA DEPENDENCY read_barrier_depends() smp_read_barrier_depends()
1619 All memory barriers except the data dependency barriers imply a compiler
1620 barrier. Data dependencies do not impose any additional compiler ordering.
1622 Aside: In the case of data dependencies, the compiler would be expected to
1623 issue the loads in the correct order (eg. `a[b]` would have to load the value
1624 of b before loading a[b]), however there is no guarantee in the C specification
1625 that the compiler may not speculate the value of b (eg. is equal to 1) and load
1626 a before b (eg. tmp = a[1]; if (b != 1) tmp = a[b]; ). There is also the
1627 problem of a compiler reloading b after having loaded a[b], thus having a newer
1628 copy of b than a[b]. A consensus has not yet been reached about these problems,
1629 however the ACCESS_ONCE macro is a good place to start looking.
1631 SMP memory barriers are reduced to compiler barriers on uniprocessor compiled
1632 systems because it is assumed that a CPU will appear to be self-consistent,
1633 and will order overlapping accesses correctly with respect to itself.
1635 [!] Note that SMP memory barriers _must_ be used to control the ordering of
1636 references to shared memory on SMP systems, though the use of locking instead
1637 is sufficient.
1639 Mandatory barriers should not be used to control SMP effects, since mandatory
1640 barriers unnecessarily impose overhead on UP systems. They may, however, be
1641 used to control MMIO effects on accesses through relaxed memory I/O windows.
1642 These are required even on non-SMP systems as they affect the order in which
1643 memory operations appear to a device by prohibiting both the compiler and the
1644 CPU from reordering them.
1647 There are some more advanced barrier functions:
1649 (*) set_mb(var, value)
1651 This assigns the value to the variable and then inserts a full memory
1652 barrier after it, depending on the function. It isn't guaranteed to
1653 insert anything more than a compiler barrier in a UP compilation.
1656 (*) smp_mb__before_atomic();
1657 (*) smp_mb__after_atomic();
1659 These are for use with atomic (such as add, subtract, increment and
1660 decrement) functions that don't return a value, especially when used for
1661 reference counting. These functions do not imply memory barriers.
1663 These are also used for atomic bitop functions that do not return a
1664 value (such as set_bit and clear_bit).
1666 As an example, consider a piece of code that marks an object as being dead
1667 and then decrements the object's reference count:
1669 obj->dead = 1;
1670 smp_mb__before_atomic();
1671 atomic_dec(&obj->ref_count);
1673 This makes sure that the death mark on the object is perceived to be set
1674 *before* the reference counter is decremented.
1676 See Documentation/atomic_ops.txt for more information. See the "Atomic
1677 operations" subsection for information on where to use these.
1680 (*) dma_wmb();
1681 (*) dma_rmb();
1683 These are for use with consistent memory to guarantee the ordering
1684 of writes or reads of shared memory accessible to both the CPU and a
1685 DMA capable device.
1687 For example, consider a device driver that shares memory with a device
1688 and uses a descriptor status value to indicate if the descriptor belongs
1689 to the device or the CPU, and a doorbell to notify it when new
1690 descriptors are available:
1692 if (desc->status != DEVICE_OWN) {
1693 /* do not read data until we own descriptor */
1694 dma_rmb();
1696 /* read/modify data */
1697 read_data = desc->data;
1698 desc->data = write_data;
1700 /* flush modifications before status update */
1701 dma_wmb();
1703 /* assign ownership */
1704 desc->status = DEVICE_OWN;
1706 /* force memory to sync before notifying device via MMIO */
1707 wmb();
1709 /* notify device of new descriptors */
1710 writel(DESC_NOTIFY, doorbell);
1711 }
1713 The dma_rmb() allows us guarantee the device has released ownership
1714 before we read the data from the descriptor, and he dma_wmb() allows
1715 us to guarantee the data is written to the descriptor before the device
1716 can see it now has ownership. The wmb() is needed to guarantee that the
1717 cache coherent memory writes have completed before attempting a write to
1718 the cache incoherent MMIO region.
1720 See Documentation/DMA-API.txt for more information on consistent memory.
1722 MMIO WRITE BARRIER
1723 ------------------
1725 The Linux kernel also has a special barrier for use with memory-mapped I/O
1726 writes:
1728 mmiowb();
1730 This is a variation on the mandatory write barrier that causes writes to weakly
1731 ordered I/O regions to be partially ordered. Its effects may go beyond the
1732 CPU->Hardware interface and actually affect the hardware at some level.
1734 See the subsection "Locks vs I/O accesses" for more information.
1737 ===============================
1738 IMPLICIT KERNEL MEMORY BARRIERS
1739 ===============================
1741 Some of the other functions in the linux kernel imply memory barriers, amongst
1742 which are locking and scheduling functions.
1744 This specification is a _minimum_ guarantee; any particular architecture may
1745 provide more substantial guarantees, but these may not be relied upon outside
1746 of arch specific code.
1749 ACQUIRING FUNCTIONS
1750 -------------------
1752 The Linux kernel has a number of locking constructs:
1754 (*) spin locks
1755 (*) R/W spin locks
1756 (*) mutexes
1757 (*) semaphores
1758 (*) R/W semaphores
1759 (*) RCU
1761 In all cases there are variants on "ACQUIRE" operations and "RELEASE" operations
1762 for each construct. These operations all imply certain barriers:
1764 (1) ACQUIRE operation implication:
1766 Memory operations issued after the ACQUIRE will be completed after the
1767 ACQUIRE operation has completed.
1769 Memory operations issued before the ACQUIRE may be completed after
1770 the ACQUIRE operation has completed. An smp_mb__before_spinlock(),
1771 combined with a following ACQUIRE, orders prior loads against
1772 subsequent loads and stores and also orders prior stores against
1773 subsequent stores. Note that this is weaker than smp_mb()! The
1774 smp_mb__before_spinlock() primitive is free on many architectures.
1776 (2) RELEASE operation implication:
1778 Memory operations issued before the RELEASE will be completed before the
1779 RELEASE operation has completed.
1781 Memory operations issued after the RELEASE may be completed before the
1782 RELEASE operation has completed.
1784 (3) ACQUIRE vs ACQUIRE implication:
1786 All ACQUIRE operations issued before another ACQUIRE operation will be
1787 completed before that ACQUIRE operation.
1789 (4) ACQUIRE vs RELEASE implication:
1791 All ACQUIRE operations issued before a RELEASE operation will be
1792 completed before the RELEASE operation.
1794 (5) Failed conditional ACQUIRE implication:
1796 Certain locking variants of the ACQUIRE operation may fail, either due to
1797 being unable to get the lock immediately, or due to receiving an unblocked
1798 signal whilst asleep waiting for the lock to become available. Failed
1799 locks do not imply any sort of barrier.
1801 [!] Note: one of the consequences of lock ACQUIREs and RELEASEs being only
1802 one-way barriers is that the effects of instructions outside of a critical
1803 section may seep into the inside of the critical section.
1805 An ACQUIRE followed by a RELEASE may not be assumed to be full memory barrier
1806 because it is possible for an access preceding the ACQUIRE to happen after the
1807 ACQUIRE, and an access following the RELEASE to happen before the RELEASE, and
1808 the two accesses can themselves then cross:
1810 *A = a;
1811 ACQUIRE M
1812 RELEASE M
1813 *B = b;
1815 may occur as:
1817 ACQUIRE M, STORE *B, STORE *A, RELEASE M
1819 When the ACQUIRE and RELEASE are a lock acquisition and release,
1820 respectively, this same reordering can occur if the lock's ACQUIRE and
1821 RELEASE are to the same lock variable, but only from the perspective of
1822 another CPU not holding that lock. In short, a ACQUIRE followed by an
1823 RELEASE may -not- be assumed to be a full memory barrier.
1825 Similarly, the reverse case of a RELEASE followed by an ACQUIRE does not
1826 imply a full memory barrier. If it is necessary for a RELEASE-ACQUIRE
1827 pair to produce a full barrier, the ACQUIRE can be followed by an
1828 smp_mb__after_unlock_lock() invocation. This will produce a full barrier
1829 if either (a) the RELEASE and the ACQUIRE are executed by the same
1830 CPU or task, or (b) the RELEASE and ACQUIRE act on the same variable.
1831 The smp_mb__after_unlock_lock() primitive is free on many architectures.
1832 Without smp_mb__after_unlock_lock(), the CPU's execution of the critical
1833 sections corresponding to the RELEASE and the ACQUIRE can cross, so that:
1835 *A = a;
1836 RELEASE M
1837 ACQUIRE N
1838 *B = b;
1840 could occur as:
1842 ACQUIRE N, STORE *B, STORE *A, RELEASE M
1844 It might appear that this reordering could introduce a deadlock.
1845 However, this cannot happen because if such a deadlock threatened,
1846 the RELEASE would simply complete, thereby avoiding the deadlock.
1848 Why does this work?
1850 One key point is that we are only talking about the CPU doing
1851 the reordering, not the compiler. If the compiler (or, for
1852 that matter, the developer) switched the operations, deadlock
1853 -could- occur.
1855 But suppose the CPU reordered the operations. In this case,
1856 the unlock precedes the lock in the assembly code. The CPU
1857 simply elected to try executing the later lock operation first.
1858 If there is a deadlock, this lock operation will simply spin (or
1859 try to sleep, but more on that later). The CPU will eventually
1860 execute the unlock operation (which preceded the lock operation
1861 in the assembly code), which will unravel the potential deadlock,
1862 allowing the lock operation to succeed.
1864 But what if the lock is a sleeplock? In that case, the code will
1865 try to enter the scheduler, where it will eventually encounter
1866 a memory barrier, which will force the earlier unlock operation
1867 to complete, again unraveling the deadlock. There might be
1868 a sleep-unlock race, but the locking primitive needs to resolve
1869 such races properly in any case.
1871 With smp_mb__after_unlock_lock(), the two critical sections cannot overlap.
1872 For example, with the following code, the store to *A will always be
1873 seen by other CPUs before the store to *B:
1875 *A = a;
1876 RELEASE M
1877 ACQUIRE N
1878 smp_mb__after_unlock_lock();
1879 *B = b;
1881 The operations will always occur in one of the following orders:
1883 STORE *A, RELEASE, ACQUIRE, smp_mb__after_unlock_lock(), STORE *B
1884 STORE *A, ACQUIRE, RELEASE, smp_mb__after_unlock_lock(), STORE *B
1885 ACQUIRE, STORE *A, RELEASE, smp_mb__after_unlock_lock(), STORE *B
1887 If the RELEASE and ACQUIRE were instead both operating on the same lock
1888 variable, only the first of these alternatives can occur. In addition,
1889 the more strongly ordered systems may rule out some of the above orders.
1890 But in any case, as noted earlier, the smp_mb__after_unlock_lock()
1891 ensures that the store to *A will always be seen as happening before
1892 the store to *B.
1894 Locks and semaphores may not provide any guarantee of ordering on UP compiled
1895 systems, and so cannot be counted on in such a situation to actually achieve
1896 anything at all - especially with respect to I/O accesses - unless combined
1897 with interrupt disabling operations.
1899 See also the section on "Inter-CPU locking barrier effects".
1902 As an example, consider the following:
1904 *A = a;
1905 *B = b;
1906 ACQUIRE
1907 *C = c;
1908 *D = d;
1909 RELEASE
1910 *E = e;
1911 *F = f;
1913 The following sequence of events is acceptable:
1915 ACQUIRE, {*F,*A}, *E, {*C,*D}, *B, RELEASE
1917 [+] Note that {*F,*A} indicates a combined access.
1919 But none of the following are:
1921 {*F,*A}, *B, ACQUIRE, *C, *D, RELEASE, *E
1922 *A, *B, *C, ACQUIRE, *D, RELEASE, *E, *F
1923 *A, *B, ACQUIRE, *C, RELEASE, *D, *E, *F
1924 *B, ACQUIRE, *C, *D, RELEASE, {*F,*A}, *E
1928 INTERRUPT DISABLING FUNCTIONS
1929 -----------------------------
1931 Functions that disable interrupts (ACQUIRE equivalent) and enable interrupts
1932 (RELEASE equivalent) will act as compiler barriers only. So if memory or I/O
1933 barriers are required in such a situation, they must be provided from some
1934 other means.
1937 SLEEP AND WAKE-UP FUNCTIONS
1938 ---------------------------
1940 Sleeping and waking on an event flagged in global data can be viewed as an
1941 interaction between two pieces of data: the task state of the task waiting for
1942 the event and the global data used to indicate the event. To make sure that
1943 these appear to happen in the right order, the primitives to begin the process
1944 of going to sleep, and the primitives to initiate a wake up imply certain
1945 barriers.
1947 Firstly, the sleeper normally follows something like this sequence of events:
1949 for (;;) {
1950 set_current_state(TASK_UNINTERRUPTIBLE);
1951 if (event_indicated)
1952 break;
1953 schedule();
1954 }
1956 A general memory barrier is interpolated automatically by set_current_state()
1957 after it has altered the task state:
1959 CPU 1
1960 ===============================
1961 set_current_state();
1962 set_mb();
1963 STORE current->state
1964 <general barrier>
1965 LOAD event_indicated
1967 set_current_state() may be wrapped by:
1969 prepare_to_wait();
1970 prepare_to_wait_exclusive();
1972 which therefore also imply a general memory barrier after setting the state.
1973 The whole sequence above is available in various canned forms, all of which
1974 interpolate the memory barrier in the right place:
1976 wait_event();
1977 wait_event_interruptible();
1978 wait_event_interruptible_exclusive();
1979 wait_event_interruptible_timeout();
1980 wait_event_killable();
1981 wait_event_timeout();
1982 wait_on_bit();
1983 wait_on_bit_lock();
1986 Secondly, code that performs a wake up normally follows something like this:
1988 event_indicated = 1;
1989 wake_up(&event_wait_queue);
1991 or:
1993 event_indicated = 1;
1994 wake_up_process(event_daemon);
1996 A write memory barrier is implied by wake_up() and co. if and only if they wake
1997 something up. The barrier occurs before the task state is cleared, and so sits
1998 between the STORE to indicate the event and the STORE to set TASK_RUNNING:
2000 CPU 1 CPU 2
2001 =============================== ===============================
2002 set_current_state(); STORE event_indicated
2003 set_mb(); wake_up();
2004 STORE current->state <write barrier>
2005 <general barrier> STORE current->state
2006 LOAD event_indicated
2008 To repeat, this write memory barrier is present if and only if something
2009 is actually awakened. To see this, consider the following sequence of
2010 events, where X and Y are both initially zero:
2012 CPU 1 CPU 2
2013 =============================== ===============================
2014 X = 1; STORE event_indicated
2015 smp_mb(); wake_up();
2016 Y = 1; wait_event(wq, Y == 1);
2017 wake_up(); load from Y sees 1, no memory barrier
2018 load from X might see 0
2020 In contrast, if a wakeup does occur, CPU 2's load from X would be guaranteed
2021 to see 1.
2023 The available waker functions include:
2025 complete();
2026 wake_up();
2027 wake_up_all();
2028 wake_up_bit();
2029 wake_up_interruptible();
2030 wake_up_interruptible_all();
2031 wake_up_interruptible_nr();
2032 wake_up_interruptible_poll();
2033 wake_up_interruptible_sync();
2034 wake_up_interruptible_sync_poll();
2035 wake_up_locked();
2036 wake_up_locked_poll();
2037 wake_up_nr();
2038 wake_up_poll();
2039 wake_up_process();
2042 [!] Note that the memory barriers implied by the sleeper and the waker do _not_
2043 order multiple stores before the wake-up with respect to loads of those stored
2044 values after the sleeper has called set_current_state(). For instance, if the
2045 sleeper does:
2047 set_current_state(TASK_INTERRUPTIBLE);
2048 if (event_indicated)
2049 break;
2050 __set_current_state(TASK_RUNNING);
2051 do_something(my_data);
2053 and the waker does:
2055 my_data = value;
2056 event_indicated = 1;
2057 wake_up(&event_wait_queue);
2059 there's no guarantee that the change to event_indicated will be perceived by
2060 the sleeper as coming after the change to my_data. In such a circumstance, the
2061 code on both sides must interpolate its own memory barriers between the
2062 separate data accesses. Thus the above sleeper ought to do:
2064 set_current_state(TASK_INTERRUPTIBLE);
2065 if (event_indicated) {
2066 smp_rmb();
2067 do_something(my_data);
2068 }
2070 and the waker should do:
2072 my_data = value;
2073 smp_wmb();
2074 event_indicated = 1;
2075 wake_up(&event_wait_queue);
2078 MISCELLANEOUS FUNCTIONS
2079 -----------------------
2081 Other functions that imply barriers:
2083 (*) schedule() and similar imply full memory barriers.
2086 ===================================
2087 INTER-CPU ACQUIRING BARRIER EFFECTS
2088 ===================================
2090 On SMP systems locking primitives give a more substantial form of barrier: one
2091 that does affect memory access ordering on other CPUs, within the context of
2092 conflict on any particular lock.
2095 ACQUIRES VS MEMORY ACCESSES
2096 ---------------------------
2098 Consider the following: the system has a pair of spinlocks (M) and (Q), and
2099 three CPUs; then should the following sequence of events occur:
2101 CPU 1 CPU 2
2102 =============================== ===============================
2103 ACCESS_ONCE(*A) = a; ACCESS_ONCE(*E) = e;
2104 ACQUIRE M ACQUIRE Q
2105 ACCESS_ONCE(*B) = b; ACCESS_ONCE(*F) = f;
2106 ACCESS_ONCE(*C) = c; ACCESS_ONCE(*G) = g;
2107 RELEASE M RELEASE Q
2108 ACCESS_ONCE(*D) = d; ACCESS_ONCE(*H) = h;
2110 Then there is no guarantee as to what order CPU 3 will see the accesses to *A
2111 through *H occur in, other than the constraints imposed by the separate locks
2112 on the separate CPUs. It might, for example, see:
2114 *E, ACQUIRE M, ACQUIRE Q, *G, *C, *F, *A, *B, RELEASE Q, *D, *H, RELEASE M
2116 But it won't see any of:
2118 *B, *C or *D preceding ACQUIRE M
2119 *A, *B or *C following RELEASE M
2120 *F, *G or *H preceding ACQUIRE Q
2121 *E, *F or *G following RELEASE Q
2124 However, if the following occurs:
2126 CPU 1 CPU 2
2127 =============================== ===============================
2128 ACCESS_ONCE(*A) = a;
2129 ACQUIRE M [1]
2130 ACCESS_ONCE(*B) = b;
2131 ACCESS_ONCE(*C) = c;
2132 RELEASE M [1]
2133 ACCESS_ONCE(*D) = d; ACCESS_ONCE(*E) = e;
2134 ACQUIRE M [2]
2135 smp_mb__after_unlock_lock();
2136 ACCESS_ONCE(*F) = f;
2137 ACCESS_ONCE(*G) = g;
2138 RELEASE M [2]
2139 ACCESS_ONCE(*H) = h;
2141 CPU 3 might see:
2143 *E, ACQUIRE M [1], *C, *B, *A, RELEASE M [1],
2144 ACQUIRE M [2], *H, *F, *G, RELEASE M [2], *D
2146 But assuming CPU 1 gets the lock first, CPU 3 won't see any of:
2148 *B, *C, *D, *F, *G or *H preceding ACQUIRE M [1]
2149 *A, *B or *C following RELEASE M [1]
2150 *F, *G or *H preceding ACQUIRE M [2]
2151 *A, *B, *C, *E, *F or *G following RELEASE M [2]
2153 Note that the smp_mb__after_unlock_lock() is critically important
2154 here: Without it CPU 3 might see some of the above orderings.
2155 Without smp_mb__after_unlock_lock(), the accesses are not guaranteed
2156 to be seen in order unless CPU 3 holds lock M.
2159 ACQUIRES VS I/O ACCESSES
2160 ------------------------
2162 Under certain circumstances (especially involving NUMA), I/O accesses within
2163 two spinlocked sections on two different CPUs may be seen as interleaved by the
2164 PCI bridge, because the PCI bridge does not necessarily participate in the
2165 cache-coherence protocol, and is therefore incapable of issuing the required
2166 read memory barriers.
2168 For example:
2170 CPU 1 CPU 2
2171 =============================== ===============================
2172 spin_lock(Q)
2173 writel(0, ADDR)
2174 writel(1, DATA);
2175 spin_unlock(Q);
2176 spin_lock(Q);
2177 writel(4, ADDR);
2178 writel(5, DATA);
2179 spin_unlock(Q);
2181 may be seen by the PCI bridge as follows:
2183 STORE *ADDR = 0, STORE *ADDR = 4, STORE *DATA = 1, STORE *DATA = 5
2185 which would probably cause the hardware to malfunction.
2188 What is necessary here is to intervene with an mmiowb() before dropping the
2189 spinlock, for example:
2191 CPU 1 CPU 2
2192 =============================== ===============================
2193 spin_lock(Q)
2194 writel(0, ADDR)
2195 writel(1, DATA);
2196 mmiowb();
2197 spin_unlock(Q);
2198 spin_lock(Q);
2199 writel(4, ADDR);
2200 writel(5, DATA);
2201 mmiowb();
2202 spin_unlock(Q);
2204 this will ensure that the two stores issued on CPU 1 appear at the PCI bridge
2205 before either of the stores issued on CPU 2.
2208 Furthermore, following a store by a load from the same device obviates the need
2209 for the mmiowb(), because the load forces the store to complete before the load
2210 is performed:
2212 CPU 1 CPU 2
2213 =============================== ===============================
2214 spin_lock(Q)
2215 writel(0, ADDR)
2216 a = readl(DATA);
2217 spin_unlock(Q);
2218 spin_lock(Q);
2219 writel(4, ADDR);
2220 b = readl(DATA);
2221 spin_unlock(Q);
2224 See Documentation/DocBook/deviceiobook.tmpl for more information.
2227 =================================
2228 WHERE ARE MEMORY BARRIERS NEEDED?
2229 =================================
2231 Under normal operation, memory operation reordering is generally not going to
2232 be a problem as a single-threaded linear piece of code will still appear to
2233 work correctly, even if it's in an SMP kernel. There are, however, four
2234 circumstances in which reordering definitely _could_ be a problem:
2236 (*) Interprocessor interaction.
2238 (*) Atomic operations.
2240 (*) Accessing devices.
2242 (*) Interrupts.
2245 INTERPROCESSOR INTERACTION
2246 --------------------------
2248 When there's a system with more than one processor, more than one CPU in the
2249 system may be working on the same data set at the same time. This can cause
2250 synchronisation problems, and the usual way of dealing with them is to use
2251 locks. Locks, however, are quite expensive, and so it may be preferable to
2252 operate without the use of a lock if at all possible. In such a case
2253 operations that affect both CPUs may have to be carefully ordered to prevent
2254 a malfunction.
2256 Consider, for example, the R/W semaphore slow path. Here a waiting process is
2257 queued on the semaphore, by virtue of it having a piece of its stack linked to
2258 the semaphore's list of waiting processes:
2260 struct rw_semaphore {
2261 ...
2262 spinlock_t lock;
2263 struct list_head waiters;
2264 };
2266 struct rwsem_waiter {
2267 struct list_head list;
2268 struct task_struct *task;
2269 };
2271 To wake up a particular waiter, the up_read() or up_write() functions have to:
2273 (1) read the next pointer from this waiter's record to know as to where the
2274 next waiter record is;
2276 (2) read the pointer to the waiter's task structure;
2278 (3) clear the task pointer to tell the waiter it has been given the semaphore;
2280 (4) call wake_up_process() on the task; and
2282 (5) release the reference held on the waiter's task struct.
2284 In other words, it has to perform this sequence of events:
2286 LOAD waiter->list.next;
2287 LOAD waiter->task;
2288 STORE waiter->task;
2289 CALL wakeup
2290 RELEASE task
2292 and if any of these steps occur out of order, then the whole thing may
2293 malfunction.
2295 Once it has queued itself and dropped the semaphore lock, the waiter does not
2296 get the lock again; it instead just waits for its task pointer to be cleared
2297 before proceeding. Since the record is on the waiter's stack, this means that
2298 if the task pointer is cleared _before_ the next pointer in the list is read,
2299 another CPU might start processing the waiter and might clobber the waiter's
2300 stack before the up*() function has a chance to read the next pointer.
2302 Consider then what might happen to the above sequence of events:
2304 CPU 1 CPU 2
2305 =============================== ===============================
2306 down_xxx()
2307 Queue waiter
2308 Sleep
2309 up_yyy()
2310 LOAD waiter->task;
2311 STORE waiter->task;
2312 Woken up by other event
2313 <preempt>
2314 Resume processing
2315 down_xxx() returns
2316 call foo()
2317 foo() clobbers *waiter
2318 </preempt>
2319 LOAD waiter->list.next;
2320 --- OOPS ---
2322 This could be dealt with using the semaphore lock, but then the down_xxx()
2323 function has to needlessly get the spinlock again after being woken up.
2325 The way to deal with this is to insert a general SMP memory barrier:
2327 LOAD waiter->list.next;
2328 LOAD waiter->task;
2329 smp_mb();
2330 STORE waiter->task;
2331 CALL wakeup
2332 RELEASE task
2334 In this case, the barrier makes a guarantee that all memory accesses before the
2335 barrier will appear to happen before all the memory accesses after the barrier
2336 with respect to the other CPUs on the system. It does _not_ guarantee that all
2337 the memory accesses before the barrier will be complete by the time the barrier
2338 instruction itself is complete.
2340 On a UP system - where this wouldn't be a problem - the smp_mb() is just a
2341 compiler barrier, thus making sure the compiler emits the instructions in the
2342 right order without actually intervening in the CPU. Since there's only one
2343 CPU, that CPU's dependency ordering logic will take care of everything else.
2346 ATOMIC OPERATIONS
2347 -----------------
2349 Whilst they are technically interprocessor interaction considerations, atomic
2350 operations are noted specially as some of them imply full memory barriers and
2351 some don't, but they're very heavily relied on as a group throughout the
2352 kernel.
2354 Any atomic operation that modifies some state in memory and returns information
2355 about the state (old or new) implies an SMP-conditional general memory barrier
2356 (smp_mb()) on each side of the actual operation (with the exception of
2357 explicit lock operations, described later). These include:
2359 xchg();
2360 cmpxchg();
2361 atomic_xchg(); atomic_long_xchg();
2362 atomic_cmpxchg(); atomic_long_cmpxchg();
2363 atomic_inc_return(); atomic_long_inc_return();
2364 atomic_dec_return(); atomic_long_dec_return();
2365 atomic_add_return(); atomic_long_add_return();
2366 atomic_sub_return(); atomic_long_sub_return();
2367 atomic_inc_and_test(); atomic_long_inc_and_test();
2368 atomic_dec_and_test(); atomic_long_dec_and_test();
2369 atomic_sub_and_test(); atomic_long_sub_and_test();
2370 atomic_add_negative(); atomic_long_add_negative();
2371 test_and_set_bit();
2372 test_and_clear_bit();
2373 test_and_change_bit();
2375 /* when succeeds (returns 1) */
2376 atomic_add_unless(); atomic_long_add_unless();
2378 These are used for such things as implementing ACQUIRE-class and RELEASE-class
2379 operations and adjusting reference counters towards object destruction, and as
2380 such the implicit memory barrier effects are necessary.
2383 The following operations are potential problems as they do _not_ imply memory
2384 barriers, but might be used for implementing such things as RELEASE-class
2385 operations:
2387 atomic_set();
2388 set_bit();
2389 clear_bit();
2390 change_bit();
2392 With these the appropriate explicit memory barrier should be used if necessary
2393 (smp_mb__before_atomic() for instance).
2396 The following also do _not_ imply memory barriers, and so may require explicit
2397 memory barriers under some circumstances (smp_mb__before_atomic() for
2398 instance):
2400 atomic_add();
2401 atomic_sub();
2402 atomic_inc();
2403 atomic_dec();
2405 If they're used for statistics generation, then they probably don't need memory
2406 barriers, unless there's a coupling between statistical data.
2408 If they're used for reference counting on an object to control its lifetime,
2409 they probably don't need memory barriers because either the reference count
2410 will be adjusted inside a locked section, or the caller will already hold
2411 sufficient references to make the lock, and thus a memory barrier unnecessary.
2413 If they're used for constructing a lock of some description, then they probably
2414 do need memory barriers as a lock primitive generally has to do things in a
2415 specific order.
2417 Basically, each usage case has to be carefully considered as to whether memory
2418 barriers are needed or not.
2420 The following operations are special locking primitives:
2422 test_and_set_bit_lock();
2423 clear_bit_unlock();
2424 __clear_bit_unlock();
2426 These implement ACQUIRE-class and RELEASE-class operations. These should be used in
2427 preference to other operations when implementing locking primitives, because
2428 their implementations can be optimised on many architectures.
2430 [!] Note that special memory barrier primitives are available for these
2431 situations because on some CPUs the atomic instructions used imply full memory
2432 barriers, and so barrier instructions are superfluous in conjunction with them,
2433 and in such cases the special barrier primitives will be no-ops.
2435 See Documentation/atomic_ops.txt for more information.
2438 ACCESSING DEVICES
2439 -----------------
2441 Many devices can be memory mapped, and so appear to the CPU as if they're just
2442 a set of memory locations. To control such a device, the driver usually has to
2443 make the right memory accesses in exactly the right order.
2445 However, having a clever CPU or a clever compiler creates a potential problem
2446 in that the carefully sequenced accesses in the driver code won't reach the
2447 device in the requisite order if the CPU or the compiler thinks it is more
2448 efficient to reorder, combine or merge accesses - something that would cause
2449 the device to malfunction.
2451 Inside of the Linux kernel, I/O should be done through the appropriate accessor
2452 routines - such as inb() or writel() - which know how to make such accesses
2453 appropriately sequential. Whilst this, for the most part, renders the explicit
2454 use of memory barriers unnecessary, there are a couple of situations where they
2455 might be needed:
2457 (1) On some systems, I/O stores are not strongly ordered across all CPUs, and
2458 so for _all_ general drivers locks should be used and mmiowb() must be
2459 issued prior to unlocking the critical section.
2461 (2) If the accessor functions are used to refer to an I/O memory window with
2462 relaxed memory access properties, then _mandatory_ memory barriers are
2463 required to enforce ordering.
2465 See Documentation/DocBook/deviceiobook.tmpl for more information.
2468 INTERRUPTS
2469 ----------
2471 A driver may be interrupted by its own interrupt service routine, and thus the
2472 two parts of the driver may interfere with each other's attempts to control or
2473 access the device.
2475 This may be alleviated - at least in part - by disabling local interrupts (a
2476 form of locking), such that the critical operations are all contained within
2477 the interrupt-disabled section in the driver. Whilst the driver's interrupt
2478 routine is executing, the driver's core may not run on the same CPU, and its
2479 interrupt is not permitted to happen again until the current interrupt has been
2480 handled, thus the interrupt handler does not need to lock against that.
2482 However, consider a driver that was talking to an ethernet card that sports an
2483 address register and a data register. If that driver's core talks to the card
2484 under interrupt-disablement and then the driver's interrupt handler is invoked:
2486 LOCAL IRQ DISABLE
2487 writew(ADDR, 3);
2488 writew(DATA, y);
2489 LOCAL IRQ ENABLE
2490 <interrupt>
2491 writew(ADDR, 4);
2492 q = readw(DATA);
2493 </interrupt>
2495 The store to the data register might happen after the second store to the
2496 address register if ordering rules are sufficiently relaxed:
2498 STORE *ADDR = 3, STORE *ADDR = 4, STORE *DATA = y, q = LOAD *DATA
2501 If ordering rules are relaxed, it must be assumed that accesses done inside an
2502 interrupt disabled section may leak outside of it and may interleave with
2503 accesses performed in an interrupt - and vice versa - unless implicit or
2504 explicit barriers are used.
2506 Normally this won't be a problem because the I/O accesses done inside such
2507 sections will include synchronous load operations on strictly ordered I/O
2508 registers that form implicit I/O barriers. If this isn't sufficient then an
2509 mmiowb() may need to be used explicitly.
2512 A similar situation may occur between an interrupt routine and two routines
2513 running on separate CPUs that communicate with each other. If such a case is
2514 likely, then interrupt-disabling locks should be used to guarantee ordering.
2517 ==========================
2518 KERNEL I/O BARRIER EFFECTS
2519 ==========================
2521 When accessing I/O memory, drivers should use the appropriate accessor
2522 functions:
2524 (*) inX(), outX():
2526 These are intended to talk to I/O space rather than memory space, but
2527 that's primarily a CPU-specific concept. The i386 and x86_64 processors do
2528 indeed have special I/O space access cycles and instructions, but many
2529 CPUs don't have such a concept.
2531 The PCI bus, amongst others, defines an I/O space concept which - on such
2532 CPUs as i386 and x86_64 - readily maps to the CPU's concept of I/O
2533 space. However, it may also be mapped as a virtual I/O space in the CPU's
2534 memory map, particularly on those CPUs that don't support alternate I/O
2535 spaces.
2537 Accesses to this space may be fully synchronous (as on i386), but
2538 intermediary bridges (such as the PCI host bridge) may not fully honour
2539 that.
2541 They are guaranteed to be fully ordered with respect to each other.
2543 They are not guaranteed to be fully ordered with respect to other types of
2544 memory and I/O operation.
2546 (*) readX(), writeX():
2548 Whether these are guaranteed to be fully ordered and uncombined with
2549 respect to each other on the issuing CPU depends on the characteristics
2550 defined for the memory window through which they're accessing. On later
2551 i386 architecture machines, for example, this is controlled by way of the
2552 MTRR registers.
2554 Ordinarily, these will be guaranteed to be fully ordered and uncombined,
2555 provided they're not accessing a prefetchable device.
2557 However, intermediary hardware (such as a PCI bridge) may indulge in
2558 deferral if it so wishes; to flush a store, a load from the same location
2559 is preferred[*], but a load from the same device or from configuration
2560 space should suffice for PCI.
2562 [*] NOTE! attempting to load from the same location as was written to may
2563 cause a malfunction - consider the 16550 Rx/Tx serial registers for
2564 example.
2566 Used with prefetchable I/O memory, an mmiowb() barrier may be required to
2567 force stores to be ordered.
2569 Please refer to the PCI specification for more information on interactions
2570 between PCI transactions.
2572 (*) readX_relaxed(), writeX_relaxed()
2574 These are similar to readX() and writeX(), but provide weaker memory
2575 ordering guarantees. Specifically, they do not guarantee ordering with
2576 respect to normal memory accesses (e.g. DMA buffers) nor do they guarantee
2577 ordering with respect to LOCK or UNLOCK operations. If the latter is
2578 required, an mmiowb() barrier can be used. Note that relaxed accesses to
2579 the same peripheral are guaranteed to be ordered with respect to each
2580 other.
2582 (*) ioreadX(), iowriteX()
2584 These will perform appropriately for the type of access they're actually
2585 doing, be it inX()/outX() or readX()/writeX().
2588 ========================================
2589 ASSUMED MINIMUM EXECUTION ORDERING MODEL
2590 ========================================
2592 It has to be assumed that the conceptual CPU is weakly-ordered but that it will
2593 maintain the appearance of program causality with respect to itself. Some CPUs
2594 (such as i386 or x86_64) are more constrained than others (such as powerpc or
2595 frv), and so the most relaxed case (namely DEC Alpha) must be assumed outside
2596 of arch-specific code.
2598 This means that it must be considered that the CPU will execute its instruction
2599 stream in any order it feels like - or even in parallel - provided that if an
2600 instruction in the stream depends on an earlier instruction, then that
2601 earlier instruction must be sufficiently complete[*] before the later
2602 instruction may proceed; in other words: provided that the appearance of
2603 causality is maintained.
2605 [*] Some instructions have more than one effect - such as changing the
2606 condition codes, changing registers or changing memory - and different
2607 instructions may depend on different effects.
2609 A CPU may also discard any instruction sequence that winds up having no
2610 ultimate effect. For example, if two adjacent instructions both load an
2611 immediate value into the same register, the first may be discarded.
2614 Similarly, it has to be assumed that compiler might reorder the instruction
2615 stream in any way it sees fit, again provided the appearance of causality is
2616 maintained.
2619 ============================
2620 THE EFFECTS OF THE CPU CACHE
2621 ============================
2623 The way cached memory operations are perceived across the system is affected to
2624 a certain extent by the caches that lie between CPUs and memory, and by the
2625 memory coherence system that maintains the consistency of state in the system.
2627 As far as the way a CPU interacts with another part of the system through the
2628 caches goes, the memory system has to include the CPU's caches, and memory
2629 barriers for the most part act at the interface between the CPU and its cache
2630 (memory barriers logically act on the dotted line in the following diagram):
2632 <--- CPU ---> : <----------- Memory ----------->
2633 :
2634 +--------+ +--------+ : +--------+ +-----------+
2635 | | | | : | | | | +--------+
2636 | CPU | | Memory | : | CPU | | | | |
2637 | Core |--->| Access |----->| Cache |<-->| | | |
2638 | | | Queue | : | | | |--->| Memory |
2639 | | | | : | | | | | |
2640 +--------+ +--------+ : +--------+ | | | |
2641 : | Cache | +--------+
2642 : | Coherency |
2643 : | Mechanism | +--------+
2644 +--------+ +--------+ : +--------+ | | | |
2645 | | | | : | | | | | |
2646 | CPU | | Memory | : | CPU | | |--->| Device |
2647 | Core |--->| Access |----->| Cache |<-->| | | |
2648 | | | Queue | : | | | | | |
2649 | | | | : | | | | +--------+
2650 +--------+ +--------+ : +--------+ +-----------+
2651 :
2652 :
2654 Although any particular load or store may not actually appear outside of the
2655 CPU that issued it since it may have been satisfied within the CPU's own cache,
2656 it will still appear as if the full memory access had taken place as far as the
2657 other CPUs are concerned since the cache coherency mechanisms will migrate the
2658 cacheline over to the accessing CPU and propagate the effects upon conflict.
2660 The CPU core may execute instructions in any order it deems fit, provided the
2661 expected program causality appears to be maintained. Some of the instructions
2662 generate load and store operations which then go into the queue of memory
2663 accesses to be performed. The core may place these in the queue in any order
2664 it wishes, and continue execution until it is forced to wait for an instruction
2665 to complete.
2667 What memory barriers are concerned with is controlling the order in which
2668 accesses cross from the CPU side of things to the memory side of things, and
2669 the order in which the effects are perceived to happen by the other observers
2670 in the system.
2672 [!] Memory barriers are _not_ needed within a given CPU, as CPUs always see
2673 their own loads and stores as if they had happened in program order.
2675 [!] MMIO or other device accesses may bypass the cache system. This depends on
2676 the properties of the memory window through which devices are accessed and/or
2677 the use of any special device communication instructions the CPU may have.
2680 CACHE COHERENCY
2681 ---------------
2683 Life isn't quite as simple as it may appear above, however: for while the
2684 caches are expected to be coherent, there's no guarantee that that coherency
2685 will be ordered. This means that whilst changes made on one CPU will
2686 eventually become visible on all CPUs, there's no guarantee that they will
2687 become apparent in the same order on those other CPUs.
2690 Consider dealing with a system that has a pair of CPUs (1 & 2), each of which
2691 has a pair of parallel data caches (CPU 1 has A/B, and CPU 2 has C/D):
2693 :
2694 : +--------+
2695 : +---------+ | |
2696 +--------+ : +--->| Cache A |<------->| |
2697 | | : | +---------+ | |
2698 | CPU 1 |<---+ | |
2699 | | : | +---------+ | |
2700 +--------+ : +--->| Cache B |<------->| |
2701 : +---------+ | |
2702 : | Memory |
2703 : +---------+ | System |
2704 +--------+ : +--->| Cache C |<------->| |
2705 | | : | +---------+ | |
2706 | CPU 2 |<---+ | |
2707 | | : | +---------+ | |
2708 +--------+ : +--->| Cache D |<------->| |
2709 : +---------+ | |
2710 : +--------+
2711 :
2713 Imagine the system has the following properties:
2715 (*) an odd-numbered cache line may be in cache A, cache C or it may still be
2716 resident in memory;
2718 (*) an even-numbered cache line may be in cache B, cache D or it may still be
2719 resident in memory;
2721 (*) whilst the CPU core is interrogating one cache, the other cache may be
2722 making use of the bus to access the rest of the system - perhaps to
2723 displace a dirty cacheline or to do a speculative load;
2725 (*) each cache has a queue of operations that need to be applied to that cache
2726 to maintain coherency with the rest of the system;
2728 (*) the coherency queue is not flushed by normal loads to lines already
2729 present in the cache, even though the contents of the queue may
2730 potentially affect those loads.
2732 Imagine, then, that two writes are made on the first CPU, with a write barrier
2733 between them to guarantee that they will appear to reach that CPU's caches in
2734 the requisite order:
2736 CPU 1 CPU 2 COMMENT
2737 =============== =============== =======================================
2738 u == 0, v == 1 and p == &u, q == &u
2739 v = 2;
2740 smp_wmb(); Make sure change to v is visible before
2741 change to p
2742 <A:modify v=2> v is now in cache A exclusively
2743 p = &v;
2744 <B:modify p=&v> p is now in cache B exclusively
2746 The write memory barrier forces the other CPUs in the system to perceive that
2747 the local CPU's caches have apparently been updated in the correct order. But
2748 now imagine that the second CPU wants to read those values:
2750 CPU 1 CPU 2 COMMENT
2751 =============== =============== =======================================
2752 ...
2753 q = p;
2754 x = *q;
2756 The above pair of reads may then fail to happen in the expected order, as the
2757 cacheline holding p may get updated in one of the second CPU's caches whilst
2758 the update to the cacheline holding v is delayed in the other of the second
2759 CPU's caches by some other cache event:
2761 CPU 1 CPU 2 COMMENT
2762 =============== =============== =======================================
2763 u == 0, v == 1 and p == &u, q == &u
2764 v = 2;
2765 smp_wmb();
2766 <A:modify v=2> <C:busy>
2767 <C:queue v=2>
2768 p = &v; q = p;
2769 <D:request p>
2770 <B:modify p=&v> <D:commit p=&v>
2771 <D:read p>
2772 x = *q;
2773 <C:read *q> Reads from v before v updated in cache
2774 <C:unbusy>
2775 <C:commit v=2>
2777 Basically, whilst both cachelines will be updated on CPU 2 eventually, there's
2778 no guarantee that, without intervention, the order of update will be the same
2779 as that committed on CPU 1.
2782 To intervene, we need to interpolate a data dependency barrier or a read
2783 barrier between the loads. This will force the cache to commit its coherency
2784 queue before processing any further requests:
2786 CPU 1 CPU 2 COMMENT
2787 =============== =============== =======================================
2788 u == 0, v == 1 and p == &u, q == &u
2789 v = 2;
2790 smp_wmb();
2791 <A:modify v=2> <C:busy>
2792 <C:queue v=2>
2793 p = &v; q = p;
2794 <D:request p>
2795 <B:modify p=&v> <D:commit p=&v>
2796 <D:read p>
2797 smp_read_barrier_depends()
2798 <C:unbusy>
2799 <C:commit v=2>
2800 x = *q;
2801 <C:read *q> Reads from v after v updated in cache
2804 This sort of problem can be encountered on DEC Alpha processors as they have a
2805 split cache that improves performance by making better use of the data bus.
2806 Whilst most CPUs do imply a data dependency barrier on the read when a memory
2807 access depends on a read, not all do, so it may not be relied on.
2809 Other CPUs may also have split caches, but must coordinate between the various
2810 cachelets for normal memory accesses. The semantics of the Alpha removes the
2811 need for coordination in the absence of memory barriers.
2814 CACHE COHERENCY VS DMA
2815 ----------------------
2817 Not all systems maintain cache coherency with respect to devices doing DMA. In
2818 such cases, a device attempting DMA may obtain stale data from RAM because
2819 dirty cache lines may be resident in the caches of various CPUs, and may not
2820 have been written back to RAM yet. To deal with this, the appropriate part of
2821 the kernel must flush the overlapping bits of cache on each CPU (and maybe
2822 invalidate them as well).
2824 In addition, the data DMA'd to RAM by a device may be overwritten by dirty
2825 cache lines being written back to RAM from a CPU's cache after the device has
2826 installed its own data, or cache lines present in the CPU's cache may simply
2827 obscure the fact that RAM has been updated, until at such time as the cacheline
2828 is discarded from the CPU's cache and reloaded. To deal with this, the
2829 appropriate part of the kernel must invalidate the overlapping bits of the
2830 cache on each CPU.
2832 See Documentation/cachetlb.txt for more information on cache management.
2835 CACHE COHERENCY VS MMIO
2836 -----------------------
2838 Memory mapped I/O usually takes place through memory locations that are part of
2839 a window in the CPU's memory space that has different properties assigned than
2840 the usual RAM directed window.
2842 Amongst these properties is usually the fact that such accesses bypass the
2843 caching entirely and go directly to the device buses. This means MMIO accesses
2844 may, in effect, overtake accesses to cached memory that were emitted earlier.
2845 A memory barrier isn't sufficient in such a case, but rather the cache must be
2846 flushed between the cached memory write and the MMIO access if the two are in
2847 any way dependent.
2850 =========================
2851 THE THINGS CPUS GET UP TO
2852 =========================
2854 A programmer might take it for granted that the CPU will perform memory
2855 operations in exactly the order specified, so that if the CPU is, for example,
2856 given the following piece of code to execute:
2858 a = ACCESS_ONCE(*A);
2859 ACCESS_ONCE(*B) = b;
2860 c = ACCESS_ONCE(*C);
2861 d = ACCESS_ONCE(*D);
2862 ACCESS_ONCE(*E) = e;
2864 they would then expect that the CPU will complete the memory operation for each
2865 instruction before moving on to the next one, leading to a definite sequence of
2866 operations as seen by external observers in the system:
2868 LOAD *A, STORE *B, LOAD *C, LOAD *D, STORE *E.
2871 Reality is, of course, much messier. With many CPUs and compilers, the above
2872 assumption doesn't hold because:
2874 (*) loads are more likely to need to be completed immediately to permit
2875 execution progress, whereas stores can often be deferred without a
2876 problem;
2878 (*) loads may be done speculatively, and the result discarded should it prove
2879 to have been unnecessary;
2881 (*) loads may be done speculatively, leading to the result having been fetched
2882 at the wrong time in the expected sequence of events;
2884 (*) the order of the memory accesses may be rearranged to promote better use
2885 of the CPU buses and caches;
2887 (*) loads and stores may be combined to improve performance when talking to
2888 memory or I/O hardware that can do batched accesses of adjacent locations,
2889 thus cutting down on transaction setup costs (memory and PCI devices may
2890 both be able to do this); and
2892 (*) the CPU's data cache may affect the ordering, and whilst cache-coherency
2893 mechanisms may alleviate this - once the store has actually hit the cache
2894 - there's no guarantee that the coherency management will be propagated in
2895 order to other CPUs.
2897 So what another CPU, say, might actually observe from the above piece of code
2898 is:
2900 LOAD *A, ..., LOAD {*C,*D}, STORE *E, STORE *B
2902 (Where "LOAD {*C,*D}" is a combined load)
2905 However, it is guaranteed that a CPU will be self-consistent: it will see its
2906 _own_ accesses appear to be correctly ordered, without the need for a memory
2907 barrier. For instance with the following code:
2909 U = ACCESS_ONCE(*A);
2910 ACCESS_ONCE(*A) = V;
2911 ACCESS_ONCE(*A) = W;
2912 X = ACCESS_ONCE(*A);
2913 ACCESS_ONCE(*A) = Y;
2914 Z = ACCESS_ONCE(*A);
2916 and assuming no intervention by an external influence, it can be assumed that
2917 the final result will appear to be:
2919 U == the original value of *A
2920 X == W
2921 Z == Y
2922 *A == Y
2924 The code above may cause the CPU to generate the full sequence of memory
2925 accesses:
2927 U=LOAD *A, STORE *A=V, STORE *A=W, X=LOAD *A, STORE *A=Y, Z=LOAD *A
2929 in that order, but, without intervention, the sequence may have almost any
2930 combination of elements combined or discarded, provided the program's view of
2931 the world remains consistent. Note that ACCESS_ONCE() is -not- optional
2932 in the above example, as there are architectures where a given CPU might
2933 reorder successive loads to the same location. On such architectures,
2934 ACCESS_ONCE() does whatever is necessary to prevent this, for example, on
2935 Itanium the volatile casts used by ACCESS_ONCE() cause GCC to emit the
2936 special ld.acq and st.rel instructions that prevent such reordering.
2938 The compiler may also combine, discard or defer elements of the sequence before
2939 the CPU even sees them.
2941 For instance:
2943 *A = V;
2944 *A = W;
2946 may be reduced to:
2948 *A = W;
2950 since, without either a write barrier or an ACCESS_ONCE(), it can be
2951 assumed that the effect of the storage of V to *A is lost. Similarly:
2953 *A = Y;
2954 Z = *A;
2956 may, without a memory barrier or an ACCESS_ONCE(), be reduced to:
2958 *A = Y;
2959 Z = Y;
2961 and the LOAD operation never appear outside of the CPU.
2964 AND THEN THERE'S THE ALPHA
2965 --------------------------
2967 The DEC Alpha CPU is one of the most relaxed CPUs there is. Not only that,
2968 some versions of the Alpha CPU have a split data cache, permitting them to have
2969 two semantically-related cache lines updated at separate times. This is where
2970 the data dependency barrier really becomes necessary as this synchronises both
2971 caches with the memory coherence system, thus making it seem like pointer
2972 changes vs new data occur in the right order.
2974 The Alpha defines the Linux kernel's memory barrier model.
2976 See the subsection on "Cache Coherency" above.
2979 ============
2980 EXAMPLE USES
2981 ============
2983 CIRCULAR BUFFERS
2984 ----------------
2986 Memory barriers can be used to implement circular buffering without the need
2987 of a lock to serialise the producer with the consumer. See:
2989 Documentation/circular-buffers.txt
2991 for details.
2994 ==========
2995 REFERENCES
2996 ==========
2998 Alpha AXP Architecture Reference Manual, Second Edition (Sites & Witek,
2999 Digital Press)
3000 Chapter 5.2: Physical Address Space Characteristics
3001 Chapter 5.4: Caches and Write Buffers
3002 Chapter 5.5: Data Sharing
3003 Chapter 5.6: Read/Write Ordering
3005 AMD64 Architecture Programmer's Manual Volume 2: System Programming
3006 Chapter 7.1: Memory-Access Ordering
3007 Chapter 7.4: Buffering and Combining Memory Writes
3009 IA-32 Intel Architecture Software Developer's Manual, Volume 3:
3010 System Programming Guide
3011 Chapter 7.1: Locked Atomic Operations
3012 Chapter 7.2: Memory Ordering
3013 Chapter 7.4: Serializing Instructions
3015 The SPARC Architecture Manual, Version 9
3016 Chapter 8: Memory Models
3017 Appendix D: Formal Specification of the Memory Models
3018 Appendix J: Programming with the Memory Models
3020 UltraSPARC Programmer Reference Manual
3021 Chapter 5: Memory Accesses and Cacheability
3022 Chapter 15: Sparc-V9 Memory Models
3024 UltraSPARC III Cu User's Manual
3025 Chapter 9: Memory Models
3027 UltraSPARC IIIi Processor User's Manual
3028 Chapter 8: Memory Models
3030 UltraSPARC Architecture 2005
3031 Chapter 9: Memory
3032 Appendix D: Formal Specifications of the Memory Models
3034 UltraSPARC T1 Supplement to the UltraSPARC Architecture 2005
3035 Chapter 8: Memory Models
3036 Appendix F: Caches and Cache Coherency
3038 Solaris Internals, Core Kernel Architecture, p63-68:
3039 Chapter 3.3: Hardware Considerations for Locks and
3040 Synchronization
3042 Unix Systems for Modern Architectures, Symmetric Multiprocessing and Caching
3043 for Kernel Programmers:
3044 Chapter 13: Other Memory Models
3046 Intel Itanium Architecture Software Developer's Manual: Volume 1:
3047 Section 2.6: Speculation
3048 Section 4.4: Memory Access