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Based on kernel version 4.8. Page generated on 2016-10-06 23:18 EST.

1				 ============================
2				 LINUX KERNEL MEMORY BARRIERS
3				 ============================
4	
5	By: David Howells <dhowells@redhat.com>
6	    Paul E. McKenney <paulmck@linux.vnet.ibm.com>
7	    Will Deacon <will.deacon@arm.com>
8	    Peter Zijlstra <peterz@infradead.org>
9	
10	==========
11	DISCLAIMER
12	==========
13	
14	This document is not a specification; it is intentionally (for the sake of
15	brevity) and unintentionally (due to being human) incomplete. This document is
16	meant as a guide to using the various memory barriers provided by Linux, but
17	in case of any doubt (and there are many) please ask.
18	
19	To repeat, this document is not a specification of what Linux expects from
20	hardware.
21	
22	The purpose of this document is twofold:
23	
24	 (1) to specify the minimum functionality that one can rely on for any
25	     particular barrier, and
26	
27	 (2) to provide a guide as to how to use the barriers that are available.
28	
29	Note that an architecture can provide more than the minimum requirement
30	for any particular barrier, but if the architecure provides less than
31	that, that architecture is incorrect.
32	
33	Note also that it is possible that a barrier may be a no-op for an
34	architecture because the way that arch works renders an explicit barrier
35	unnecessary in that case.
36	
37	
38	========
39	CONTENTS
40	========
41	
42	 (*) Abstract memory access model.
43	
44	     - Device operations.
45	     - Guarantees.
46	
47	 (*) What are memory barriers?
48	
49	     - Varieties of memory barrier.
50	     - What may not be assumed about memory barriers?
51	     - Data dependency barriers.
52	     - Control dependencies.
53	     - SMP barrier pairing.
54	     - Examples of memory barrier sequences.
55	     - Read memory barriers vs load speculation.
56	     - Transitivity
57	
58	 (*) Explicit kernel barriers.
59	
60	     - Compiler barrier.
61	     - CPU memory barriers.
62	     - MMIO write barrier.
63	
64	 (*) Implicit kernel memory barriers.
65	
66	     - Lock acquisition functions.
67	     - Interrupt disabling functions.
68	     - Sleep and wake-up functions.
69	     - Miscellaneous functions.
70	
71	 (*) Inter-CPU acquiring barrier effects.
72	
73	     - Acquires vs memory accesses.
74	     - Acquires vs I/O accesses.
75	
76	 (*) Where are memory barriers needed?
77	
78	     - Interprocessor interaction.
79	     - Atomic operations.
80	     - Accessing devices.
81	     - Interrupts.
82	
83	 (*) Kernel I/O barrier effects.
84	
85	 (*) Assumed minimum execution ordering model.
86	
87	 (*) The effects of the cpu cache.
88	
89	     - Cache coherency.
90	     - Cache coherency vs DMA.
91	     - Cache coherency vs MMIO.
92	
93	 (*) The things CPUs get up to.
94	
95	     - And then there's the Alpha.
96	     - Virtual Machine Guests.
97	
98	 (*) Example uses.
99	
100	     - Circular buffers.
101	
102	 (*) References.
103	
104	
105	============================
106	ABSTRACT MEMORY ACCESS MODEL
107	============================
108	
109	Consider the following abstract model of the system:
110	
111			            :                :
112			            :                :
113			            :                :
114			+-------+   :   +--------+   :   +-------+
115			|       |   :   |        |   :   |       |
116			|       |   :   |        |   :   |       |
117			| CPU 1 |<----->| Memory |<----->| CPU 2 |
118			|       |   :   |        |   :   |       |
119			|       |   :   |        |   :   |       |
120			+-------+   :   +--------+   :   +-------+
121			    ^       :       ^        :       ^
122			    |       :       |        :       |
123			    |       :       |        :       |
124			    |       :       v        :       |
125			    |       :   +--------+   :       |
126			    |       :   |        |   :       |
127			    |       :   |        |   :       |
128			    +---------->| Device |<----------+
129			            :   |        |   :
130			            :   |        |   :
131			            :   +--------+   :
132			            :                :
133	
134	Each CPU executes a program that generates memory access operations.  In the
135	abstract CPU, memory operation ordering is very relaxed, and a CPU may actually
136	perform the memory operations in any order it likes, provided program causality
137	appears to be maintained.  Similarly, the compiler may also arrange the
138	instructions it emits in any order it likes, provided it doesn't affect the
139	apparent operation of the program.
140	
141	So in the above diagram, the effects of the memory operations performed by a
142	CPU are perceived by the rest of the system as the operations cross the
143	interface between the CPU and rest of the system (the dotted lines).
144	
145	
146	For example, consider the following sequence of events:
147	
148		CPU 1		CPU 2
149		===============	===============
150		{ A == 1; B == 2 }
151		A = 3;		x = B;
152		B = 4;		y = A;
153	
154	The set of accesses as seen by the memory system in the middle can be arranged
155	in 24 different combinations:
156	
157		STORE A=3,	STORE B=4,	y=LOAD A->3,	x=LOAD B->4
158		STORE A=3,	STORE B=4,	x=LOAD B->4,	y=LOAD A->3
159		STORE A=3,	y=LOAD A->3,	STORE B=4,	x=LOAD B->4
160		STORE A=3,	y=LOAD A->3,	x=LOAD B->2,	STORE B=4
161		STORE A=3,	x=LOAD B->2,	STORE B=4,	y=LOAD A->3
162		STORE A=3,	x=LOAD B->2,	y=LOAD A->3,	STORE B=4
163		STORE B=4,	STORE A=3,	y=LOAD A->3,	x=LOAD B->4
164		STORE B=4, ...
165		...
166	
167	and can thus result in four different combinations of values:
168	
169		x == 2, y == 1
170		x == 2, y == 3
171		x == 4, y == 1
172		x == 4, y == 3
173	
174	
175	Furthermore, the stores committed by a CPU to the memory system may not be
176	perceived by the loads made by another CPU in the same order as the stores were
177	committed.
178	
179	
180	As a further example, consider this sequence of events:
181	
182		CPU 1		CPU 2
183		===============	===============
184		{ A == 1, B == 2, C == 3, P == &A, Q == &C }
185		B = 4;		Q = P;
186		P = &B		D = *Q;
187	
188	There is an obvious data dependency here, as the value loaded into D depends on
189	the address retrieved from P by CPU 2.  At the end of the sequence, any of the
190	following results are possible:
191	
192		(Q == &A) and (D == 1)
193		(Q == &B) and (D == 2)
194		(Q == &B) and (D == 4)
195	
196	Note that CPU 2 will never try and load C into D because the CPU will load P
197	into Q before issuing the load of *Q.
198	
199	
200	DEVICE OPERATIONS
201	-----------------
202	
203	Some devices present their control interfaces as collections of memory
204	locations, but the order in which the control registers are accessed is very
205	important.  For instance, imagine an ethernet card with a set of internal
206	registers that are accessed through an address port register (A) and a data
207	port register (D).  To read internal register 5, the following code might then
208	be used:
209	
210		*A = 5;
211		x = *D;
212	
213	but this might show up as either of the following two sequences:
214	
215		STORE *A = 5, x = LOAD *D
216		x = LOAD *D, STORE *A = 5
217	
218	the second of which will almost certainly result in a malfunction, since it set
219	the address _after_ attempting to read the register.
220	
221	
222	GUARANTEES
223	----------
224	
225	There are some minimal guarantees that may be expected of a CPU:
226	
227	 (*) On any given CPU, dependent memory accesses will be issued in order, with
228	     respect to itself.  This means that for:
229	
230		Q = READ_ONCE(P); smp_read_barrier_depends(); D = READ_ONCE(*Q);
231	
232	     the CPU will issue the following memory operations:
233	
234		Q = LOAD P, D = LOAD *Q
235	
236	     and always in that order.  On most systems, smp_read_barrier_depends()
237	     does nothing, but it is required for DEC Alpha.  The READ_ONCE()
238	     is required to prevent compiler mischief.  Please note that you
239	     should normally use something like rcu_dereference() instead of
240	     open-coding smp_read_barrier_depends().
241	
242	 (*) Overlapping loads and stores within a particular CPU will appear to be
243	     ordered within that CPU.  This means that for:
244	
245		a = READ_ONCE(*X); WRITE_ONCE(*X, b);
246	
247	     the CPU will only issue the following sequence of memory operations:
248	
249		a = LOAD *X, STORE *X = b
250	
251	     And for:
252	
253		WRITE_ONCE(*X, c); d = READ_ONCE(*X);
254	
255	     the CPU will only issue:
256	
257		STORE *X = c, d = LOAD *X
258	
259	     (Loads and stores overlap if they are targeted at overlapping pieces of
260	     memory).
261	
262	And there are a number of things that _must_ or _must_not_ be assumed:
263	
264	 (*) It _must_not_ be assumed that the compiler will do what you want
265	     with memory references that are not protected by READ_ONCE() and
266	     WRITE_ONCE().  Without them, the compiler is within its rights to
267	     do all sorts of "creative" transformations, which are covered in
268	     the COMPILER BARRIER section.
269	
270	 (*) It _must_not_ be assumed that independent loads and stores will be issued
271	     in the order given.  This means that for:
272	
273		X = *A; Y = *B; *D = Z;
274	
275	     we may get any of the following sequences:
276	
277		X = LOAD *A,  Y = LOAD *B,  STORE *D = Z
278		X = LOAD *A,  STORE *D = Z, Y = LOAD *B
279		Y = LOAD *B,  X = LOAD *A,  STORE *D = Z
280		Y = LOAD *B,  STORE *D = Z, X = LOAD *A
281		STORE *D = Z, X = LOAD *A,  Y = LOAD *B
282		STORE *D = Z, Y = LOAD *B,  X = LOAD *A
283	
284	 (*) It _must_ be assumed that overlapping memory accesses may be merged or
285	     discarded.  This means that for:
286	
287		X = *A; Y = *(A + 4);
288	
289	     we may get any one of the following sequences:
290	
291		X = LOAD *A; Y = LOAD *(A + 4);
292		Y = LOAD *(A + 4); X = LOAD *A;
293		{X, Y} = LOAD {*A, *(A + 4) };
294	
295	     And for:
296	
297		*A = X; *(A + 4) = Y;
298	
299	     we may get any of:
300	
301		STORE *A = X; STORE *(A + 4) = Y;
302		STORE *(A + 4) = Y; STORE *A = X;
303		STORE {*A, *(A + 4) } = {X, Y};
304	
305	And there are anti-guarantees:
306	
307	 (*) These guarantees do not apply to bitfields, because compilers often
308	     generate code to modify these using non-atomic read-modify-write
309	     sequences.  Do not attempt to use bitfields to synchronize parallel
310	     algorithms.
311	
312	 (*) Even in cases where bitfields are protected by locks, all fields
313	     in a given bitfield must be protected by one lock.  If two fields
314	     in a given bitfield are protected by different locks, the compiler's
315	     non-atomic read-modify-write sequences can cause an update to one
316	     field to corrupt the value of an adjacent field.
317	
318	 (*) These guarantees apply only to properly aligned and sized scalar
319	     variables.  "Properly sized" currently means variables that are
320	     the same size as "char", "short", "int" and "long".  "Properly
321	     aligned" means the natural alignment, thus no constraints for
322	     "char", two-byte alignment for "short", four-byte alignment for
323	     "int", and either four-byte or eight-byte alignment for "long",
324	     on 32-bit and 64-bit systems, respectively.  Note that these
325	     guarantees were introduced into the C11 standard, so beware when
326	     using older pre-C11 compilers (for example, gcc 4.6).  The portion
327	     of the standard containing this guarantee is Section 3.14, which
328	     defines "memory location" as follows:
329	
330	     	memory location
331			either an object of scalar type, or a maximal sequence
332			of adjacent bit-fields all having nonzero width
333	
334			NOTE 1: Two threads of execution can update and access
335			separate memory locations without interfering with
336			each other.
337	
338			NOTE 2: A bit-field and an adjacent non-bit-field member
339			are in separate memory locations. The same applies
340			to two bit-fields, if one is declared inside a nested
341			structure declaration and the other is not, or if the two
342			are separated by a zero-length bit-field declaration,
343			or if they are separated by a non-bit-field member
344			declaration. It is not safe to concurrently update two
345			bit-fields in the same structure if all members declared
346			between them are also bit-fields, no matter what the
347			sizes of those intervening bit-fields happen to be.
348	
349	
350	=========================
351	WHAT ARE MEMORY BARRIERS?
352	=========================
353	
354	As can be seen above, independent memory operations are effectively performed
355	in random order, but this can be a problem for CPU-CPU interaction and for I/O.
356	What is required is some way of intervening to instruct the compiler and the
357	CPU to restrict the order.
358	
359	Memory barriers are such interventions.  They impose a perceived partial
360	ordering over the memory operations on either side of the barrier.
361	
362	Such enforcement is important because the CPUs and other devices in a system
363	can use a variety of tricks to improve performance, including reordering,
364	deferral and combination of memory operations; speculative loads; speculative
365	branch prediction and various types of caching.  Memory barriers are used to
366	override or suppress these tricks, allowing the code to sanely control the
367	interaction of multiple CPUs and/or devices.
368	
369	
370	VARIETIES OF MEMORY BARRIER
371	---------------------------
372	
373	Memory barriers come in four basic varieties:
374	
375	 (1) Write (or store) memory barriers.
376	
377	     A write memory barrier gives a guarantee that all the STORE operations
378	     specified before the barrier will appear to happen before all the STORE
379	     operations specified after the barrier with respect to the other
380	     components of the system.
381	
382	     A write barrier is a partial ordering on stores only; it is not required
383	     to have any effect on loads.
384	
385	     A CPU can be viewed as committing a sequence of store operations to the
386	     memory system as time progresses.  All stores before a write barrier will
387	     occur in the sequence _before_ all the stores after the write barrier.
388	
389	     [!] Note that write barriers should normally be paired with read or data
390	     dependency barriers; see the "SMP barrier pairing" subsection.
391	
392	
393	 (2) Data dependency barriers.
394	
395	     A data dependency barrier is a weaker form of read barrier.  In the case
396	     where two loads are performed such that the second depends on the result
397	     of the first (eg: the first load retrieves the address to which the second
398	     load will be directed), a data dependency barrier would be required to
399	     make sure that the target of the second load is updated before the address
400	     obtained by the first load is accessed.
401	
402	     A data dependency barrier is a partial ordering on interdependent loads
403	     only; it is not required to have any effect on stores, independent loads
404	     or overlapping loads.
405	
406	     As mentioned in (1), the other CPUs in the system can be viewed as
407	     committing sequences of stores to the memory system that the CPU being
408	     considered can then perceive.  A data dependency barrier issued by the CPU
409	     under consideration guarantees that for any load preceding it, if that
410	     load touches one of a sequence of stores from another CPU, then by the
411	     time the barrier completes, the effects of all the stores prior to that
412	     touched by the load will be perceptible to any loads issued after the data
413	     dependency barrier.
414	
415	     See the "Examples of memory barrier sequences" subsection for diagrams
416	     showing the ordering constraints.
417	
418	     [!] Note that the first load really has to have a _data_ dependency and
419	     not a control dependency.  If the address for the second load is dependent
420	     on the first load, but the dependency is through a conditional rather than
421	     actually loading the address itself, then it's a _control_ dependency and
422	     a full read barrier or better is required.  See the "Control dependencies"
423	     subsection for more information.
424	
425	     [!] Note that data dependency barriers should normally be paired with
426	     write barriers; see the "SMP barrier pairing" subsection.
427	
428	
429	 (3) Read (or load) memory barriers.
430	
431	     A read barrier is a data dependency barrier plus a guarantee that all the
432	     LOAD operations specified before the barrier will appear to happen before
433	     all the LOAD operations specified after the barrier with respect to the
434	     other components of the system.
435	
436	     A read barrier is a partial ordering on loads only; it is not required to
437	     have any effect on stores.
438	
439	     Read memory barriers imply data dependency barriers, and so can substitute
440	     for them.
441	
442	     [!] Note that read barriers should normally be paired with write barriers;
443	     see the "SMP barrier pairing" subsection.
444	
445	
446	 (4) General memory barriers.
447	
448	     A general memory barrier gives a guarantee that all the LOAD and STORE
449	     operations specified before the barrier will appear to happen before all
450	     the LOAD and STORE operations specified after the barrier with respect to
451	     the other components of the system.
452	
453	     A general memory barrier is a partial ordering over both loads and stores.
454	
455	     General memory barriers imply both read and write memory barriers, and so
456	     can substitute for either.
457	
458	
459	And a couple of implicit varieties:
460	
461	 (5) ACQUIRE operations.
462	
463	     This acts as a one-way permeable barrier.  It guarantees that all memory
464	     operations after the ACQUIRE operation will appear to happen after the
465	     ACQUIRE operation with respect to the other components of the system.
466	     ACQUIRE operations include LOCK operations and both smp_load_acquire()
467	     and smp_cond_acquire() operations. The later builds the necessary ACQUIRE
468	     semantics from relying on a control dependency and smp_rmb().
469	
470	     Memory operations that occur before an ACQUIRE operation may appear to
471	     happen after it completes.
472	
473	     An ACQUIRE operation should almost always be paired with a RELEASE
474	     operation.
475	
476	
477	 (6) RELEASE operations.
478	
479	     This also acts as a one-way permeable barrier.  It guarantees that all
480	     memory operations before the RELEASE operation will appear to happen
481	     before the RELEASE operation with respect to the other components of the
482	     system. RELEASE operations include UNLOCK operations and
483	     smp_store_release() operations.
484	
485	     Memory operations that occur after a RELEASE operation may appear to
486	     happen before it completes.
487	
488	     The use of ACQUIRE and RELEASE operations generally precludes the need
489	     for other sorts of memory barrier (but note the exceptions mentioned in
490	     the subsection "MMIO write barrier").  In addition, a RELEASE+ACQUIRE
491	     pair is -not- guaranteed to act as a full memory barrier.  However, after
492	     an ACQUIRE on a given variable, all memory accesses preceding any prior
493	     RELEASE on that same variable are guaranteed to be visible.  In other
494	     words, within a given variable's critical section, all accesses of all
495	     previous critical sections for that variable are guaranteed to have
496	     completed.
497	
498	     This means that ACQUIRE acts as a minimal "acquire" operation and
499	     RELEASE acts as a minimal "release" operation.
500	
501	A subset of the atomic operations described in atomic_ops.txt have ACQUIRE
502	and RELEASE variants in addition to fully-ordered and relaxed (no barrier
503	semantics) definitions.  For compound atomics performing both a load and a
504	store, ACQUIRE semantics apply only to the load and RELEASE semantics apply
505	only to the store portion of the operation.
506	
507	Memory barriers are only required where there's a possibility of interaction
508	between two CPUs or between a CPU and a device.  If it can be guaranteed that
509	there won't be any such interaction in any particular piece of code, then
510	memory barriers are unnecessary in that piece of code.
511	
512	
513	Note that these are the _minimum_ guarantees.  Different architectures may give
514	more substantial guarantees, but they may _not_ be relied upon outside of arch
515	specific code.
516	
517	
518	WHAT MAY NOT BE ASSUMED ABOUT MEMORY BARRIERS?
519	----------------------------------------------
520	
521	There are certain things that the Linux kernel memory barriers do not guarantee:
522	
523	 (*) There is no guarantee that any of the memory accesses specified before a
524	     memory barrier will be _complete_ by the completion of a memory barrier
525	     instruction; the barrier can be considered to draw a line in that CPU's
526	     access queue that accesses of the appropriate type may not cross.
527	
528	 (*) There is no guarantee that issuing a memory barrier on one CPU will have
529	     any direct effect on another CPU or any other hardware in the system.  The
530	     indirect effect will be the order in which the second CPU sees the effects
531	     of the first CPU's accesses occur, but see the next point:
532	
533	 (*) There is no guarantee that a CPU will see the correct order of effects
534	     from a second CPU's accesses, even _if_ the second CPU uses a memory
535	     barrier, unless the first CPU _also_ uses a matching memory barrier (see
536	     the subsection on "SMP Barrier Pairing").
537	
538	 (*) There is no guarantee that some intervening piece of off-the-CPU
539	     hardware[*] will not reorder the memory accesses.  CPU cache coherency
540	     mechanisms should propagate the indirect effects of a memory barrier
541	     between CPUs, but might not do so in order.
542	
543		[*] For information on bus mastering DMA and coherency please read:
544	
545		    Documentation/PCI/pci.txt
546		    Documentation/DMA-API-HOWTO.txt
547		    Documentation/DMA-API.txt
548	
549	
550	DATA DEPENDENCY BARRIERS
551	------------------------
552	
553	The usage requirements of data dependency barriers are a little subtle, and
554	it's not always obvious that they're needed.  To illustrate, consider the
555	following sequence of events:
556	
557		CPU 1		      CPU 2
558		===============	      ===============
559		{ A == 1, B == 2, C == 3, P == &A, Q == &C }
560		B = 4;
561		<write barrier>
562		WRITE_ONCE(P, &B)
563				      Q = READ_ONCE(P);
564				      D = *Q;
565	
566	There's a clear data dependency here, and it would seem that by the end of the
567	sequence, Q must be either &A or &B, and that:
568	
569		(Q == &A) implies (D == 1)
570		(Q == &B) implies (D == 4)
571	
572	But!  CPU 2's perception of P may be updated _before_ its perception of B, thus
573	leading to the following situation:
574	
575		(Q == &B) and (D == 2) ????
576	
577	Whilst this may seem like a failure of coherency or causality maintenance, it
578	isn't, and this behaviour can be observed on certain real CPUs (such as the DEC
579	Alpha).
580	
581	To deal with this, a data dependency barrier or better must be inserted
582	between the address load and the data load:
583	
584		CPU 1		      CPU 2
585		===============	      ===============
586		{ A == 1, B == 2, C == 3, P == &A, Q == &C }
587		B = 4;
588		<write barrier>
589		WRITE_ONCE(P, &B);
590				      Q = READ_ONCE(P);
591				      <data dependency barrier>
592				      D = *Q;
593	
594	This enforces the occurrence of one of the two implications, and prevents the
595	third possibility from arising.
596	
597	A data-dependency barrier must also order against dependent writes:
598	
599		CPU 1		      CPU 2
600		===============	      ===============
601		{ A == 1, B == 2, C = 3, P == &A, Q == &C }
602		B = 4;
603		<write barrier>
604		WRITE_ONCE(P, &B);
605				      Q = READ_ONCE(P);
606				      <data dependency barrier>
607				      *Q = 5;
608	
609	The data-dependency barrier must order the read into Q with the store
610	into *Q.  This prohibits this outcome:
611	
612		(Q == B) && (B == 4)
613	
614	Please note that this pattern should be rare.  After all, the whole point
615	of dependency ordering is to -prevent- writes to the data structure, along
616	with the expensive cache misses associated with those writes.  This pattern
617	can be used to record rare error conditions and the like, and the ordering
618	prevents such records from being lost.
619	
620	
621	[!] Note that this extremely counterintuitive situation arises most easily on
622	machines with split caches, so that, for example, one cache bank processes
623	even-numbered cache lines and the other bank processes odd-numbered cache
624	lines.  The pointer P might be stored in an odd-numbered cache line, and the
625	variable B might be stored in an even-numbered cache line.  Then, if the
626	even-numbered bank of the reading CPU's cache is extremely busy while the
627	odd-numbered bank is idle, one can see the new value of the pointer P (&B),
628	but the old value of the variable B (2).
629	
630	
631	The data dependency barrier is very important to the RCU system,
632	for example.  See rcu_assign_pointer() and rcu_dereference() in
633	include/linux/rcupdate.h.  This permits the current target of an RCU'd
634	pointer to be replaced with a new modified target, without the replacement
635	target appearing to be incompletely initialised.
636	
637	See also the subsection on "Cache Coherency" for a more thorough example.
638	
639	
640	CONTROL DEPENDENCIES
641	--------------------
642	
643	A load-load control dependency requires a full read memory barrier, not
644	simply a data dependency barrier to make it work correctly.  Consider the
645	following bit of code:
646	
647		q = READ_ONCE(a);
648		if (q) {
649			<data dependency barrier>  /* BUG: No data dependency!!! */
650			p = READ_ONCE(b);
651		}
652	
653	This will not have the desired effect because there is no actual data
654	dependency, but rather a control dependency that the CPU may short-circuit
655	by attempting to predict the outcome in advance, so that other CPUs see
656	the load from b as having happened before the load from a.  In such a
657	case what's actually required is:
658	
659		q = READ_ONCE(a);
660		if (q) {
661			<read barrier>
662			p = READ_ONCE(b);
663		}
664	
665	However, stores are not speculated.  This means that ordering -is- provided
666	for load-store control dependencies, as in the following example:
667	
668		q = READ_ONCE(a);
669		if (q) {
670			WRITE_ONCE(b, p);
671		}
672	
673	Control dependencies pair normally with other types of barriers.  That
674	said, please note that READ_ONCE() is not optional! Without the
675	READ_ONCE(), the compiler might combine the load from 'a' with other
676	loads from 'a', and the store to 'b' with other stores to 'b', with
677	possible highly counterintuitive effects on ordering.
678	
679	Worse yet, if the compiler is able to prove (say) that the value of
680	variable 'a' is always non-zero, it would be well within its rights
681	to optimize the original example by eliminating the "if" statement
682	as follows:
683	
684		q = a;
685		b = p;  /* BUG: Compiler and CPU can both reorder!!! */
686	
687	So don't leave out the READ_ONCE().
688	
689	It is tempting to try to enforce ordering on identical stores on both
690	branches of the "if" statement as follows:
691	
692		q = READ_ONCE(a);
693		if (q) {
694			barrier();
695			WRITE_ONCE(b, p);
696			do_something();
697		} else {
698			barrier();
699			WRITE_ONCE(b, p);
700			do_something_else();
701		}
702	
703	Unfortunately, current compilers will transform this as follows at high
704	optimization levels:
705	
706		q = READ_ONCE(a);
707		barrier();
708		WRITE_ONCE(b, p);  /* BUG: No ordering vs. load from a!!! */
709		if (q) {
710			/* WRITE_ONCE(b, p); -- moved up, BUG!!! */
711			do_something();
712		} else {
713			/* WRITE_ONCE(b, p); -- moved up, BUG!!! */
714			do_something_else();
715		}
716	
717	Now there is no conditional between the load from 'a' and the store to
718	'b', which means that the CPU is within its rights to reorder them:
719	The conditional is absolutely required, and must be present in the
720	assembly code even after all compiler optimizations have been applied.
721	Therefore, if you need ordering in this example, you need explicit
722	memory barriers, for example, smp_store_release():
723	
724		q = READ_ONCE(a);
725		if (q) {
726			smp_store_release(&b, p);
727			do_something();
728		} else {
729			smp_store_release(&b, p);
730			do_something_else();
731		}
732	
733	In contrast, without explicit memory barriers, two-legged-if control
734	ordering is guaranteed only when the stores differ, for example:
735	
736		q = READ_ONCE(a);
737		if (q) {
738			WRITE_ONCE(b, p);
739			do_something();
740		} else {
741			WRITE_ONCE(b, r);
742			do_something_else();
743		}
744	
745	The initial READ_ONCE() is still required to prevent the compiler from
746	proving the value of 'a'.
747	
748	In addition, you need to be careful what you do with the local variable 'q',
749	otherwise the compiler might be able to guess the value and again remove
750	the needed conditional.  For example:
751	
752		q = READ_ONCE(a);
753		if (q % MAX) {
754			WRITE_ONCE(b, p);
755			do_something();
756		} else {
757			WRITE_ONCE(b, r);
758			do_something_else();
759		}
760	
761	If MAX is defined to be 1, then the compiler knows that (q % MAX) is
762	equal to zero, in which case the compiler is within its rights to
763	transform the above code into the following:
764	
765		q = READ_ONCE(a);
766		WRITE_ONCE(b, p);
767		do_something_else();
768	
769	Given this transformation, the CPU is not required to respect the ordering
770	between the load from variable 'a' and the store to variable 'b'.  It is
771	tempting to add a barrier(), but this does not help.  The conditional
772	is gone, and the barrier won't bring it back.  Therefore, if you are
773	relying on this ordering, you should make sure that MAX is greater than
774	one, perhaps as follows:
775	
776		q = READ_ONCE(a);
777		BUILD_BUG_ON(MAX <= 1); /* Order load from a with store to b. */
778		if (q % MAX) {
779			WRITE_ONCE(b, p);
780			do_something();
781		} else {
782			WRITE_ONCE(b, r);
783			do_something_else();
784		}
785	
786	Please note once again that the stores to 'b' differ.  If they were
787	identical, as noted earlier, the compiler could pull this store outside
788	of the 'if' statement.
789	
790	You must also be careful not to rely too much on boolean short-circuit
791	evaluation.  Consider this example:
792	
793		q = READ_ONCE(a);
794		if (q || 1 > 0)
795			WRITE_ONCE(b, 1);
796	
797	Because the first condition cannot fault and the second condition is
798	always true, the compiler can transform this example as following,
799	defeating control dependency:
800	
801		q = READ_ONCE(a);
802		WRITE_ONCE(b, 1);
803	
804	This example underscores the need to ensure that the compiler cannot
805	out-guess your code.  More generally, although READ_ONCE() does force
806	the compiler to actually emit code for a given load, it does not force
807	the compiler to use the results.
808	
809	In addition, control dependencies apply only to the then-clause and
810	else-clause of the if-statement in question.  In particular, it does
811	not necessarily apply to code following the if-statement:
812	
813		q = READ_ONCE(a);
814		if (q) {
815			WRITE_ONCE(b, p);
816		} else {
817			WRITE_ONCE(b, r);
818		}
819		WRITE_ONCE(c, 1);  /* BUG: No ordering against the read from "a". */
820	
821	It is tempting to argue that there in fact is ordering because the
822	compiler cannot reorder volatile accesses and also cannot reorder
823	the writes to "b" with the condition.  Unfortunately for this line
824	of reasoning, the compiler might compile the two writes to "b" as
825	conditional-move instructions, as in this fanciful pseudo-assembly
826	language:
827	
828		ld r1,a
829		ld r2,p
830		ld r3,r
831		cmp r1,$0
832		cmov,ne r4,r2
833		cmov,eq r4,r3
834		st r4,b
835		st $1,c
836	
837	A weakly ordered CPU would have no dependency of any sort between the load
838	from "a" and the store to "c".  The control dependencies would extend
839	only to the pair of cmov instructions and the store depending on them.
840	In short, control dependencies apply only to the stores in the then-clause
841	and else-clause of the if-statement in question (including functions
842	invoked by those two clauses), not to code following that if-statement.
843	
844	Finally, control dependencies do -not- provide transitivity.  This is
845	demonstrated by two related examples, with the initial values of
846	x and y both being zero:
847	
848		CPU 0                     CPU 1
849		=======================   =======================
850		r1 = READ_ONCE(x);        r2 = READ_ONCE(y);
851		if (r1 > 0)               if (r2 > 0)
852		  WRITE_ONCE(y, 1);         WRITE_ONCE(x, 1);
853	
854		assert(!(r1 == 1 && r2 == 1));
855	
856	The above two-CPU example will never trigger the assert().  However,
857	if control dependencies guaranteed transitivity (which they do not),
858	then adding the following CPU would guarantee a related assertion:
859	
860		CPU 2
861		=====================
862		WRITE_ONCE(x, 2);
863	
864		assert(!(r1 == 2 && r2 == 1 && x == 2)); /* FAILS!!! */
865	
866	But because control dependencies do -not- provide transitivity, the above
867	assertion can fail after the combined three-CPU example completes.  If you
868	need the three-CPU example to provide ordering, you will need smp_mb()
869	between the loads and stores in the CPU 0 and CPU 1 code fragments,
870	that is, just before or just after the "if" statements.  Furthermore,
871	the original two-CPU example is very fragile and should be avoided.
872	
873	These two examples are the LB and WWC litmus tests from this paper:
874	http://www.cl.cam.ac.uk/users/pes20/ppc-supplemental/test6.pdf and this
875	site: https://www.cl.cam.ac.uk/~pes20/ppcmem/index.html.
876	
877	In summary:
878	
879	  (*) Control dependencies can order prior loads against later stores.
880	      However, they do -not- guarantee any other sort of ordering:
881	      Not prior loads against later loads, nor prior stores against
882	      later anything.  If you need these other forms of ordering,
883	      use smp_rmb(), smp_wmb(), or, in the case of prior stores and
884	      later loads, smp_mb().
885	
886	  (*) If both legs of the "if" statement begin with identical stores to
887	      the same variable, then those stores must be ordered, either by
888	      preceding both of them with smp_mb() or by using smp_store_release()
889	      to carry out the stores.  Please note that it is -not- sufficient
890	      to use barrier() at beginning of each leg of the "if" statement
891	      because, as shown by the example above, optimizing compilers can
892	      destroy the control dependency while respecting the letter of the
893	      barrier() law.
894	
895	  (*) Control dependencies require at least one run-time conditional
896	      between the prior load and the subsequent store, and this
897	      conditional must involve the prior load.  If the compiler is able
898	      to optimize the conditional away, it will have also optimized
899	      away the ordering.  Careful use of READ_ONCE() and WRITE_ONCE()
900	      can help to preserve the needed conditional.
901	
902	  (*) Control dependencies require that the compiler avoid reordering the
903	      dependency into nonexistence.  Careful use of READ_ONCE() or
904	      atomic{,64}_read() can help to preserve your control dependency.
905	      Please see the COMPILER BARRIER section for more information.
906	
907	  (*) Control dependencies apply only to the then-clause and else-clause
908	      of the if-statement containing the control dependency, including
909	      any functions that these two clauses call.  Control dependencies
910	      do -not- apply to code following the if-statement containing the
911	      control dependency.
912	
913	  (*) Control dependencies pair normally with other types of barriers.
914	
915	  (*) Control dependencies do -not- provide transitivity.  If you
916	      need transitivity, use smp_mb().
917	
918	
919	SMP BARRIER PAIRING
920	-------------------
921	
922	When dealing with CPU-CPU interactions, certain types of memory barrier should
923	always be paired.  A lack of appropriate pairing is almost certainly an error.
924	
925	General barriers pair with each other, though they also pair with most
926	other types of barriers, albeit without transitivity.  An acquire barrier
927	pairs with a release barrier, but both may also pair with other barriers,
928	including of course general barriers.  A write barrier pairs with a data
929	dependency barrier, a control dependency, an acquire barrier, a release
930	barrier, a read barrier, or a general barrier.  Similarly a read barrier,
931	control dependency, or a data dependency barrier pairs with a write
932	barrier, an acquire barrier, a release barrier, or a general barrier:
933	
934		CPU 1		      CPU 2
935		===============	      ===============
936		WRITE_ONCE(a, 1);
937		<write barrier>
938		WRITE_ONCE(b, 2);     x = READ_ONCE(b);
939				      <read barrier>
940				      y = READ_ONCE(a);
941	
942	Or:
943	
944		CPU 1		      CPU 2
945		===============	      ===============================
946		a = 1;
947		<write barrier>
948		WRITE_ONCE(b, &a);    x = READ_ONCE(b);
949				      <data dependency barrier>
950				      y = *x;
951	
952	Or even:
953	
954		CPU 1		      CPU 2
955		===============	      ===============================
956		r1 = READ_ONCE(y);
957		<general barrier>
958		WRITE_ONCE(y, 1);     if (r2 = READ_ONCE(x)) {
959				         <implicit control dependency>
960				         WRITE_ONCE(y, 1);
961				      }
962	
963		assert(r1 == 0 || r2 == 0);
964	
965	Basically, the read barrier always has to be there, even though it can be of
966	the "weaker" type.
967	
968	[!] Note that the stores before the write barrier would normally be expected to
969	match the loads after the read barrier or the data dependency barrier, and vice
970	versa:
971	
972		CPU 1                               CPU 2
973		===================                 ===================
974		WRITE_ONCE(a, 1);    }----   --->{  v = READ_ONCE(c);
975		WRITE_ONCE(b, 2);    }    \ /    {  w = READ_ONCE(d);
976		<write barrier>            \        <read barrier>
977		WRITE_ONCE(c, 3);    }    / \    {  x = READ_ONCE(a);
978		WRITE_ONCE(d, 4);    }----   --->{  y = READ_ONCE(b);
979	
980	
981	EXAMPLES OF MEMORY BARRIER SEQUENCES
982	------------------------------------
983	
984	Firstly, write barriers act as partial orderings on store operations.
985	Consider the following sequence of events:
986	
987		CPU 1
988		=======================
989		STORE A = 1
990		STORE B = 2
991		STORE C = 3
992		<write barrier>
993		STORE D = 4
994		STORE E = 5
995	
996	This sequence of events is committed to the memory coherence system in an order
997	that the rest of the system might perceive as the unordered set of { STORE A,
998	STORE B, STORE C } all occurring before the unordered set of { STORE D, STORE E
999	}:
1000	
1001		+-------+       :      :
1002		|       |       +------+
1003		|       |------>| C=3  |     }     /\
1004		|       |  :    +------+     }-----  \  -----> Events perceptible to
1005		|       |  :    | A=1  |     }        \/       the rest of the system
1006		|       |  :    +------+     }
1007		| CPU 1 |  :    | B=2  |     }
1008		|       |       +------+     }
1009		|       |   wwwwwwwwwwwwwwww }   <--- At this point the write barrier
1010		|       |       +------+     }        requires all stores prior to the
1011		|       |  :    | E=5  |     }        barrier to be committed before
1012		|       |  :    +------+     }        further stores may take place
1013		|       |------>| D=4  |     }
1014		|       |       +------+
1015		+-------+       :      :
1016		                   |
1017		                   | Sequence in which stores are committed to the
1018		                   | memory system by CPU 1
1019		                   V
1020	
1021	
1022	Secondly, data dependency barriers act as partial orderings on data-dependent
1023	loads.  Consider the following sequence of events:
1024	
1025		CPU 1			CPU 2
1026		=======================	=======================
1027			{ B = 7; X = 9; Y = 8; C = &Y }
1028		STORE A = 1
1029		STORE B = 2
1030		<write barrier>
1031		STORE C = &B		LOAD X
1032		STORE D = 4		LOAD C (gets &B)
1033					LOAD *C (reads B)
1034	
1035	Without intervention, CPU 2 may perceive the events on CPU 1 in some
1036	effectively random order, despite the write barrier issued by CPU 1:
1037	
1038		+-------+       :      :                :       :
1039		|       |       +------+                +-------+  | Sequence of update
1040		|       |------>| B=2  |-----       --->| Y->8  |  | of perception on
1041		|       |  :    +------+     \          +-------+  | CPU 2
1042		| CPU 1 |  :    | A=1  |      \     --->| C->&Y |  V
1043		|       |       +------+       |        +-------+
1044		|       |   wwwwwwwwwwwwwwww   |        :       :
1045		|       |       +------+       |        :       :
1046		|       |  :    | C=&B |---    |        :       :       +-------+
1047		|       |  :    +------+   \   |        +-------+       |       |
1048		|       |------>| D=4  |    ----------->| C->&B |------>|       |
1049		|       |       +------+       |        +-------+       |       |
1050		+-------+       :      :       |        :       :       |       |
1051		                               |        :       :       |       |
1052		                               |        :       :       | CPU 2 |
1053		                               |        +-------+       |       |
1054		    Apparently incorrect --->  |        | B->7  |------>|       |
1055		    perception of B (!)        |        +-------+       |       |
1056		                               |        :       :       |       |
1057		                               |        +-------+       |       |
1058		    The load of X holds --->    \       | X->9  |------>|       |
1059		    up the maintenance           \      +-------+       |       |
1060		    of coherence of B             ----->| B->2  |       +-------+
1061		                                        +-------+
1062		                                        :       :
1063	
1064	
1065	In the above example, CPU 2 perceives that B is 7, despite the load of *C
1066	(which would be B) coming after the LOAD of C.
1067	
1068	If, however, a data dependency barrier were to be placed between the load of C
1069	and the load of *C (ie: B) on CPU 2:
1070	
1071		CPU 1			CPU 2
1072		=======================	=======================
1073			{ B = 7; X = 9; Y = 8; C = &Y }
1074		STORE A = 1
1075		STORE B = 2
1076		<write barrier>
1077		STORE C = &B		LOAD X
1078		STORE D = 4		LOAD C (gets &B)
1079					<data dependency barrier>
1080					LOAD *C (reads B)
1081	
1082	then the following will occur:
1083	
1084		+-------+       :      :                :       :
1085		|       |       +------+                +-------+
1086		|       |------>| B=2  |-----       --->| Y->8  |
1087		|       |  :    +------+     \          +-------+
1088		| CPU 1 |  :    | A=1  |      \     --->| C->&Y |
1089		|       |       +------+       |        +-------+
1090		|       |   wwwwwwwwwwwwwwww   |        :       :
1091		|       |       +------+       |        :       :
1092		|       |  :    | C=&B |---    |        :       :       +-------+
1093		|       |  :    +------+   \   |        +-------+       |       |
1094		|       |------>| D=4  |    ----------->| C->&B |------>|       |
1095		|       |       +------+       |        +-------+       |       |
1096		+-------+       :      :       |        :       :       |       |
1097		                               |        :       :       |       |
1098		                               |        :       :       | CPU 2 |
1099		                               |        +-------+       |       |
1100		                               |        | X->9  |------>|       |
1101		                               |        +-------+       |       |
1102		  Makes sure all effects --->   \   ddddddddddddddddd   |       |
1103		  prior to the store of C        \      +-------+       |       |
1104		  are perceptible to              ----->| B->2  |------>|       |
1105		  subsequent loads                      +-------+       |       |
1106		                                        :       :       +-------+
1107	
1108	
1109	And thirdly, a read barrier acts as a partial order on loads.  Consider the
1110	following sequence of events:
1111	
1112		CPU 1			CPU 2
1113		=======================	=======================
1114			{ A = 0, B = 9 }
1115		STORE A=1
1116		<write barrier>
1117		STORE B=2
1118					LOAD B
1119					LOAD A
1120	
1121	Without intervention, CPU 2 may then choose to perceive the events on CPU 1 in
1122	some effectively random order, despite the write barrier issued by CPU 1:
1123	
1124		+-------+       :      :                :       :
1125		|       |       +------+                +-------+
1126		|       |------>| A=1  |------      --->| A->0  |
1127		|       |       +------+      \         +-------+
1128		| CPU 1 |   wwwwwwwwwwwwwwww   \    --->| B->9  |
1129		|       |       +------+        |       +-------+
1130		|       |------>| B=2  |---     |       :       :
1131		|       |       +------+   \    |       :       :       +-------+
1132		+-------+       :      :    \   |       +-------+       |       |
1133		                             ---------->| B->2  |------>|       |
1134		                                |       +-------+       | CPU 2 |
1135		                                |       | A->0  |------>|       |
1136		                                |       +-------+       |       |
1137		                                |       :       :       +-------+
1138		                                 \      :       :
1139		                                  \     +-------+
1140		                                   ---->| A->1  |
1141		                                        +-------+
1142		                                        :       :
1143	
1144	
1145	If, however, a read barrier were to be placed between the load of B and the
1146	load of A on CPU 2:
1147	
1148		CPU 1			CPU 2
1149		=======================	=======================
1150			{ A = 0, B = 9 }
1151		STORE A=1
1152		<write barrier>
1153		STORE B=2
1154					LOAD B
1155					<read barrier>
1156					LOAD A
1157	
1158	then the partial ordering imposed by CPU 1 will be perceived correctly by CPU
1159	2:
1160	
1161		+-------+       :      :                :       :
1162		|       |       +------+                +-------+
1163		|       |------>| A=1  |------      --->| A->0  |
1164		|       |       +------+      \         +-------+
1165		| CPU 1 |   wwwwwwwwwwwwwwww   \    --->| B->9  |
1166		|       |       +------+        |       +-------+
1167		|       |------>| B=2  |---     |       :       :
1168		|       |       +------+   \    |       :       :       +-------+
1169		+-------+       :      :    \   |       +-------+       |       |
1170		                             ---------->| B->2  |------>|       |
1171		                                |       +-------+       | CPU 2 |
1172		                                |       :       :       |       |
1173		                                |       :       :       |       |
1174		  At this point the read ---->   \  rrrrrrrrrrrrrrrrr   |       |
1175		  barrier causes all effects      \     +-------+       |       |
1176		  prior to the storage of B        ---->| A->1  |------>|       |
1177		  to be perceptible to CPU 2            +-------+       |       |
1178		                                        :       :       +-------+
1179	
1180	
1181	To illustrate this more completely, consider what could happen if the code
1182	contained a load of A either side of the read barrier:
1183	
1184		CPU 1			CPU 2
1185		=======================	=======================
1186			{ A = 0, B = 9 }
1187		STORE A=1
1188		<write barrier>
1189		STORE B=2
1190					LOAD B
1191					LOAD A [first load of A]
1192					<read barrier>
1193					LOAD A [second load of A]
1194	
1195	Even though the two loads of A both occur after the load of B, they may both
1196	come up with different values:
1197	
1198		+-------+       :      :                :       :
1199		|       |       +------+                +-------+
1200		|       |------>| A=1  |------      --->| A->0  |
1201		|       |       +------+      \         +-------+
1202		| CPU 1 |   wwwwwwwwwwwwwwww   \    --->| B->9  |
1203		|       |       +------+        |       +-------+
1204		|       |------>| B=2  |---     |       :       :
1205		|       |       +------+   \    |       :       :       +-------+
1206		+-------+       :      :    \   |       +-------+       |       |
1207		                             ---------->| B->2  |------>|       |
1208		                                |       +-------+       | CPU 2 |
1209		                                |       :       :       |       |
1210		                                |       :       :       |       |
1211		                                |       +-------+       |       |
1212		                                |       | A->0  |------>| 1st   |
1213		                                |       +-------+       |       |
1214		  At this point the read ---->   \  rrrrrrrrrrrrrrrrr   |       |
1215		  barrier causes all effects      \     +-------+       |       |
1216		  prior to the storage of B        ---->| A->1  |------>| 2nd   |
1217		  to be perceptible to CPU 2            +-------+       |       |
1218		                                        :       :       +-------+
1219	
1220	
1221	But it may be that the update to A from CPU 1 becomes perceptible to CPU 2
1222	before the read barrier completes anyway:
1223	
1224		+-------+       :      :                :       :
1225		|       |       +------+                +-------+
1226		|       |------>| A=1  |------      --->| A->0  |
1227		|       |       +------+      \         +-------+
1228		| CPU 1 |   wwwwwwwwwwwwwwww   \    --->| B->9  |
1229		|       |       +------+        |       +-------+
1230		|       |------>| B=2  |---     |       :       :
1231		|       |       +------+   \    |       :       :       +-------+
1232		+-------+       :      :    \   |       +-------+       |       |
1233		                             ---------->| B->2  |------>|       |
1234		                                |       +-------+       | CPU 2 |
1235		                                |       :       :       |       |
1236		                                 \      :       :       |       |
1237		                                  \     +-------+       |       |
1238		                                   ---->| A->1  |------>| 1st   |
1239		                                        +-------+       |       |
1240		                                    rrrrrrrrrrrrrrrrr   |       |
1241		                                        +-------+       |       |
1242		                                        | A->1  |------>| 2nd   |
1243		                                        +-------+       |       |
1244		                                        :       :       +-------+
1245	
1246	
1247	The guarantee is that the second load will always come up with A == 1 if the
1248	load of B came up with B == 2.  No such guarantee exists for the first load of
1249	A; that may come up with either A == 0 or A == 1.
1250	
1251	
1252	READ MEMORY BARRIERS VS LOAD SPECULATION
1253	----------------------------------------
1254	
1255	Many CPUs speculate with loads: that is they see that they will need to load an
1256	item from memory, and they find a time where they're not using the bus for any
1257	other loads, and so do the load in advance - even though they haven't actually
1258	got to that point in the instruction execution flow yet.  This permits the
1259	actual load instruction to potentially complete immediately because the CPU
1260	already has the value to hand.
1261	
1262	It may turn out that the CPU didn't actually need the value - perhaps because a
1263	branch circumvented the load - in which case it can discard the value or just
1264	cache it for later use.
1265	
1266	Consider:
1267	
1268		CPU 1			CPU 2
1269		=======================	=======================
1270					LOAD B
1271					DIVIDE		} Divide instructions generally
1272					DIVIDE		} take a long time to perform
1273					LOAD A
1274	
1275	Which might appear as this:
1276	
1277		                                        :       :       +-------+
1278		                                        +-------+       |       |
1279		                                    --->| B->2  |------>|       |
1280		                                        +-------+       | CPU 2 |
1281		                                        :       :DIVIDE |       |
1282		                                        +-------+       |       |
1283		The CPU being busy doing a --->     --->| A->0  |~~~~   |       |
1284		division speculates on the              +-------+   ~   |       |
1285		LOAD of A                               :       :   ~   |       |
1286		                                        :       :DIVIDE |       |
1287		                                        :       :   ~   |       |
1288		Once the divisions are complete -->     :       :   ~-->|       |
1289		the CPU can then perform the            :       :       |       |
1290		LOAD with immediate effect              :       :       +-------+
1291	
1292	
1293	Placing a read barrier or a data dependency barrier just before the second
1294	load:
1295	
1296		CPU 1			CPU 2
1297		=======================	=======================
1298					LOAD B
1299					DIVIDE
1300					DIVIDE
1301					<read barrier>
1302					LOAD A
1303	
1304	will force any value speculatively obtained to be reconsidered to an extent
1305	dependent on the type of barrier used.  If there was no change made to the
1306	speculated memory location, then the speculated value will just be used:
1307	
1308		                                        :       :       +-------+
1309		                                        +-------+       |       |
1310		                                    --->| B->2  |------>|       |
1311		                                        +-------+       | CPU 2 |
1312		                                        :       :DIVIDE |       |
1313		                                        +-------+       |       |
1314		The CPU being busy doing a --->     --->| A->0  |~~~~   |       |
1315		division speculates on the              +-------+   ~   |       |
1316		LOAD of A                               :       :   ~   |       |
1317		                                        :       :DIVIDE |       |
1318		                                        :       :   ~   |       |
1319		                                        :       :   ~   |       |
1320		                                    rrrrrrrrrrrrrrrr~   |       |
1321		                                        :       :   ~   |       |
1322		                                        :       :   ~-->|       |
1323		                                        :       :       |       |
1324		                                        :       :       +-------+
1325	
1326	
1327	but if there was an update or an invalidation from another CPU pending, then
1328	the speculation will be cancelled and the value reloaded:
1329	
1330		                                        :       :       +-------+
1331		                                        +-------+       |       |
1332		                                    --->| B->2  |------>|       |
1333		                                        +-------+       | CPU 2 |
1334		                                        :       :DIVIDE |       |
1335		                                        +-------+       |       |
1336		The CPU being busy doing a --->     --->| A->0  |~~~~   |       |
1337		division speculates on the              +-------+   ~   |       |
1338		LOAD of A                               :       :   ~   |       |
1339		                                        :       :DIVIDE |       |
1340		                                        :       :   ~   |       |
1341		                                        :       :   ~   |       |
1342		                                    rrrrrrrrrrrrrrrrr   |       |
1343		                                        +-------+       |       |
1344		The speculation is discarded --->   --->| A->1  |------>|       |
1345		and an updated value is                 +-------+       |       |
1346		retrieved                               :       :       +-------+
1347	
1348	
1349	TRANSITIVITY
1350	------------
1351	
1352	Transitivity is a deeply intuitive notion about ordering that is not
1353	always provided by real computer systems.  The following example
1354	demonstrates transitivity:
1355	
1356		CPU 1			CPU 2			CPU 3
1357		=======================	=======================	=======================
1358			{ X = 0, Y = 0 }
1359		STORE X=1		LOAD X			STORE Y=1
1360					<general barrier>	<general barrier>
1361					LOAD Y			LOAD X
1362	
1363	Suppose that CPU 2's load from X returns 1 and its load from Y returns 0.
1364	This indicates that CPU 2's load from X in some sense follows CPU 1's
1365	store to X and that CPU 2's load from Y in some sense preceded CPU 3's
1366	store to Y.  The question is then "Can CPU 3's load from X return 0?"
1367	
1368	Because CPU 2's load from X in some sense came after CPU 1's store, it
1369	is natural to expect that CPU 3's load from X must therefore return 1.
1370	This expectation is an example of transitivity: if a load executing on
1371	CPU A follows a load from the same variable executing on CPU B, then
1372	CPU A's load must either return the same value that CPU B's load did,
1373	or must return some later value.
1374	
1375	In the Linux kernel, use of general memory barriers guarantees
1376	transitivity.  Therefore, in the above example, if CPU 2's load from X
1377	returns 1 and its load from Y returns 0, then CPU 3's load from X must
1378	also return 1.
1379	
1380	However, transitivity is -not- guaranteed for read or write barriers.
1381	For example, suppose that CPU 2's general barrier in the above example
1382	is changed to a read barrier as shown below:
1383	
1384		CPU 1			CPU 2			CPU 3
1385		=======================	=======================	=======================
1386			{ X = 0, Y = 0 }
1387		STORE X=1		LOAD X			STORE Y=1
1388					<read barrier>		<general barrier>
1389					LOAD Y			LOAD X
1390	
1391	This substitution destroys transitivity: in this example, it is perfectly
1392	legal for CPU 2's load from X to return 1, its load from Y to return 0,
1393	and CPU 3's load from X to return 0.
1394	
1395	The key point is that although CPU 2's read barrier orders its pair
1396	of loads, it does not guarantee to order CPU 1's store.  Therefore, if
1397	this example runs on a system where CPUs 1 and 2 share a store buffer
1398	or a level of cache, CPU 2 might have early access to CPU 1's writes.
1399	General barriers are therefore required to ensure that all CPUs agree
1400	on the combined order of CPU 1's and CPU 2's accesses.
1401	
1402	General barriers provide "global transitivity", so that all CPUs will
1403	agree on the order of operations.  In contrast, a chain of release-acquire
1404	pairs provides only "local transitivity", so that only those CPUs on
1405	the chain are guaranteed to agree on the combined order of the accesses.
1406	For example, switching to C code in deference to Herman Hollerith:
1407	
1408		int u, v, x, y, z;
1409	
1410		void cpu0(void)
1411		{
1412			r0 = smp_load_acquire(&x);
1413			WRITE_ONCE(u, 1);
1414			smp_store_release(&y, 1);
1415		}
1416	
1417		void cpu1(void)
1418		{
1419			r1 = smp_load_acquire(&y);
1420			r4 = READ_ONCE(v);
1421			r5 = READ_ONCE(u);
1422			smp_store_release(&z, 1);
1423		}
1424	
1425		void cpu2(void)
1426		{
1427			r2 = smp_load_acquire(&z);
1428			smp_store_release(&x, 1);
1429		}
1430	
1431		void cpu3(void)
1432		{
1433			WRITE_ONCE(v, 1);
1434			smp_mb();
1435			r3 = READ_ONCE(u);
1436		}
1437	
1438	Because cpu0(), cpu1(), and cpu2() participate in a local transitive
1439	chain of smp_store_release()/smp_load_acquire() pairs, the following
1440	outcome is prohibited:
1441	
1442		r0 == 1 && r1 == 1 && r2 == 1
1443	
1444	Furthermore, because of the release-acquire relationship between cpu0()
1445	and cpu1(), cpu1() must see cpu0()'s writes, so that the following
1446	outcome is prohibited:
1447	
1448		r1 == 1 && r5 == 0
1449	
1450	However, the transitivity of release-acquire is local to the participating
1451	CPUs and does not apply to cpu3().  Therefore, the following outcome
1452	is possible:
1453	
1454		r0 == 0 && r1 == 1 && r2 == 1 && r3 == 0 && r4 == 0
1455	
1456	As an aside, the following outcome is also possible:
1457	
1458		r0 == 0 && r1 == 1 && r2 == 1 && r3 == 0 && r4 == 0 && r5 == 1
1459	
1460	Although cpu0(), cpu1(), and cpu2() will see their respective reads and
1461	writes in order, CPUs not involved in the release-acquire chain might
1462	well disagree on the order.  This disagreement stems from the fact that
1463	the weak memory-barrier instructions used to implement smp_load_acquire()
1464	and smp_store_release() are not required to order prior stores against
1465	subsequent loads in all cases.  This means that cpu3() can see cpu0()'s
1466	store to u as happening -after- cpu1()'s load from v, even though
1467	both cpu0() and cpu1() agree that these two operations occurred in the
1468	intended order.
1469	
1470	However, please keep in mind that smp_load_acquire() is not magic.
1471	In particular, it simply reads from its argument with ordering.  It does
1472	-not- ensure that any particular value will be read.  Therefore, the
1473	following outcome is possible:
1474	
1475		r0 == 0 && r1 == 0 && r2 == 0 && r5 == 0
1476	
1477	Note that this outcome can happen even on a mythical sequentially
1478	consistent system where nothing is ever reordered.
1479	
1480	To reiterate, if your code requires global transitivity, use general
1481	barriers throughout.
1482	
1483	
1484	========================
1485	EXPLICIT KERNEL BARRIERS
1486	========================
1487	
1488	The Linux kernel has a variety of different barriers that act at different
1489	levels:
1490	
1491	  (*) Compiler barrier.
1492	
1493	  (*) CPU memory barriers.
1494	
1495	  (*) MMIO write barrier.
1496	
1497	
1498	COMPILER BARRIER
1499	----------------
1500	
1501	The Linux kernel has an explicit compiler barrier function that prevents the
1502	compiler from moving the memory accesses either side of it to the other side:
1503	
1504		barrier();
1505	
1506	This is a general barrier -- there are no read-read or write-write
1507	variants of barrier().  However, READ_ONCE() and WRITE_ONCE() can be
1508	thought of as weak forms of barrier() that affect only the specific
1509	accesses flagged by the READ_ONCE() or WRITE_ONCE().
1510	
1511	The barrier() function has the following effects:
1512	
1513	 (*) Prevents the compiler from reordering accesses following the
1514	     barrier() to precede any accesses preceding the barrier().
1515	     One example use for this property is to ease communication between
1516	     interrupt-handler code and the code that was interrupted.
1517	
1518	 (*) Within a loop, forces the compiler to load the variables used
1519	     in that loop's conditional on each pass through that loop.
1520	
1521	The READ_ONCE() and WRITE_ONCE() functions can prevent any number of
1522	optimizations that, while perfectly safe in single-threaded code, can
1523	be fatal in concurrent code.  Here are some examples of these sorts
1524	of optimizations:
1525	
1526	 (*) The compiler is within its rights to reorder loads and stores
1527	     to the same variable, and in some cases, the CPU is within its
1528	     rights to reorder loads to the same variable.  This means that
1529	     the following code:
1530	
1531		a[0] = x;
1532		a[1] = x;
1533	
1534	     Might result in an older value of x stored in a[1] than in a[0].
1535	     Prevent both the compiler and the CPU from doing this as follows:
1536	
1537		a[0] = READ_ONCE(x);
1538		a[1] = READ_ONCE(x);
1539	
1540	     In short, READ_ONCE() and WRITE_ONCE() provide cache coherence for
1541	     accesses from multiple CPUs to a single variable.
1542	
1543	 (*) The compiler is within its rights to merge successive loads from
1544	     the same variable.  Such merging can cause the compiler to "optimize"
1545	     the following code:
1546	
1547		while (tmp = a)
1548			do_something_with(tmp);
1549	
1550	     into the following code, which, although in some sense legitimate
1551	     for single-threaded code, is almost certainly not what the developer
1552	     intended:
1553	
1554		if (tmp = a)
1555			for (;;)
1556				do_something_with(tmp);
1557	
1558	     Use READ_ONCE() to prevent the compiler from doing this to you:
1559	
1560		while (tmp = READ_ONCE(a))
1561			do_something_with(tmp);
1562	
1563	 (*) The compiler is within its rights to reload a variable, for example,
1564	     in cases where high register pressure prevents the compiler from
1565	     keeping all data of interest in registers.  The compiler might
1566	     therefore optimize the variable 'tmp' out of our previous example:
1567	
1568		while (tmp = a)
1569			do_something_with(tmp);
1570	
1571	     This could result in the following code, which is perfectly safe in
1572	     single-threaded code, but can be fatal in concurrent code:
1573	
1574		while (a)
1575			do_something_with(a);
1576	
1577	     For example, the optimized version of this code could result in
1578	     passing a zero to do_something_with() in the case where the variable
1579	     a was modified by some other CPU between the "while" statement and
1580	     the call to do_something_with().
1581	
1582	     Again, use READ_ONCE() to prevent the compiler from doing this:
1583	
1584		while (tmp = READ_ONCE(a))
1585			do_something_with(tmp);
1586	
1587	     Note that if the compiler runs short of registers, it might save
1588	     tmp onto the stack.  The overhead of this saving and later restoring
1589	     is why compilers reload variables.  Doing so is perfectly safe for
1590	     single-threaded code, so you need to tell the compiler about cases
1591	     where it is not safe.
1592	
1593	 (*) The compiler is within its rights to omit a load entirely if it knows
1594	     what the value will be.  For example, if the compiler can prove that
1595	     the value of variable 'a' is always zero, it can optimize this code:
1596	
1597		while (tmp = a)
1598			do_something_with(tmp);
1599	
1600	     Into this:
1601	
1602		do { } while (0);
1603	
1604	     This transformation is a win for single-threaded code because it
1605	     gets rid of a load and a branch.  The problem is that the compiler
1606	     will carry out its proof assuming that the current CPU is the only
1607	     one updating variable 'a'.  If variable 'a' is shared, then the
1608	     compiler's proof will be erroneous.  Use READ_ONCE() to tell the
1609	     compiler that it doesn't know as much as it thinks it does:
1610	
1611		while (tmp = READ_ONCE(a))
1612			do_something_with(tmp);
1613	
1614	     But please note that the compiler is also closely watching what you
1615	     do with the value after the READ_ONCE().  For example, suppose you
1616	     do the following and MAX is a preprocessor macro with the value 1:
1617	
1618		while ((tmp = READ_ONCE(a)) % MAX)
1619			do_something_with(tmp);
1620	
1621	     Then the compiler knows that the result of the "%" operator applied
1622	     to MAX will always be zero, again allowing the compiler to optimize
1623	     the code into near-nonexistence.  (It will still load from the
1624	     variable 'a'.)
1625	
1626	 (*) Similarly, the compiler is within its rights to omit a store entirely
1627	     if it knows that the variable already has the value being stored.
1628	     Again, the compiler assumes that the current CPU is the only one
1629	     storing into the variable, which can cause the compiler to do the
1630	     wrong thing for shared variables.  For example, suppose you have
1631	     the following:
1632	
1633		a = 0;
1634		... Code that does not store to variable a ...
1635		a = 0;
1636	
1637	     The compiler sees that the value of variable 'a' is already zero, so
1638	     it might well omit the second store.  This would come as a fatal
1639	     surprise if some other CPU might have stored to variable 'a' in the
1640	     meantime.
1641	
1642	     Use WRITE_ONCE() to prevent the compiler from making this sort of
1643	     wrong guess:
1644	
1645		WRITE_ONCE(a, 0);
1646		... Code that does not store to variable a ...
1647		WRITE_ONCE(a, 0);
1648	
1649	 (*) The compiler is within its rights to reorder memory accesses unless
1650	     you tell it not to.  For example, consider the following interaction
1651	     between process-level code and an interrupt handler:
1652	
1653		void process_level(void)
1654		{
1655			msg = get_message();
1656			flag = true;
1657		}
1658	
1659		void interrupt_handler(void)
1660		{
1661			if (flag)
1662				process_message(msg);
1663		}
1664	
1665	     There is nothing to prevent the compiler from transforming
1666	     process_level() to the following, in fact, this might well be a
1667	     win for single-threaded code:
1668	
1669		void process_level(void)
1670		{
1671			flag = true;
1672			msg = get_message();
1673		}
1674	
1675	     If the interrupt occurs between these two statement, then
1676	     interrupt_handler() might be passed a garbled msg.  Use WRITE_ONCE()
1677	     to prevent this as follows:
1678	
1679		void process_level(void)
1680		{
1681			WRITE_ONCE(msg, get_message());
1682			WRITE_ONCE(flag, true);
1683		}
1684	
1685		void interrupt_handler(void)
1686		{
1687			if (READ_ONCE(flag))
1688				process_message(READ_ONCE(msg));
1689		}
1690	
1691	     Note that the READ_ONCE() and WRITE_ONCE() wrappers in
1692	     interrupt_handler() are needed if this interrupt handler can itself
1693	     be interrupted by something that also accesses 'flag' and 'msg',
1694	     for example, a nested interrupt or an NMI.  Otherwise, READ_ONCE()
1695	     and WRITE_ONCE() are not needed in interrupt_handler() other than
1696	     for documentation purposes.  (Note also that nested interrupts
1697	     do not typically occur in modern Linux kernels, in fact, if an
1698	     interrupt handler returns with interrupts enabled, you will get a
1699	     WARN_ONCE() splat.)
1700	
1701	     You should assume that the compiler can move READ_ONCE() and
1702	     WRITE_ONCE() past code not containing READ_ONCE(), WRITE_ONCE(),
1703	     barrier(), or similar primitives.
1704	
1705	     This effect could also be achieved using barrier(), but READ_ONCE()
1706	     and WRITE_ONCE() are more selective:  With READ_ONCE() and
1707	     WRITE_ONCE(), the compiler need only forget the contents of the
1708	     indicated memory locations, while with barrier() the compiler must
1709	     discard the value of all memory locations that it has currented
1710	     cached in any machine registers.  Of course, the compiler must also
1711	     respect the order in which the READ_ONCE()s and WRITE_ONCE()s occur,
1712	     though the CPU of course need not do so.
1713	
1714	 (*) The compiler is within its rights to invent stores to a variable,
1715	     as in the following example:
1716	
1717		if (a)
1718			b = a;
1719		else
1720			b = 42;
1721	
1722	     The compiler might save a branch by optimizing this as follows:
1723	
1724		b = 42;
1725		if (a)
1726			b = a;
1727	
1728	     In single-threaded code, this is not only safe, but also saves
1729	     a branch.  Unfortunately, in concurrent code, this optimization
1730	     could cause some other CPU to see a spurious value of 42 -- even
1731	     if variable 'a' was never zero -- when loading variable 'b'.
1732	     Use WRITE_ONCE() to prevent this as follows:
1733	
1734		if (a)
1735			WRITE_ONCE(b, a);
1736		else
1737			WRITE_ONCE(b, 42);
1738	
1739	     The compiler can also invent loads.  These are usually less
1740	     damaging, but they can result in cache-line bouncing and thus in
1741	     poor performance and scalability.  Use READ_ONCE() to prevent
1742	     invented loads.
1743	
1744	 (*) For aligned memory locations whose size allows them to be accessed
1745	     with a single memory-reference instruction, prevents "load tearing"
1746	     and "store tearing," in which a single large access is replaced by
1747	     multiple smaller accesses.  For example, given an architecture having
1748	     16-bit store instructions with 7-bit immediate fields, the compiler
1749	     might be tempted to use two 16-bit store-immediate instructions to
1750	     implement the following 32-bit store:
1751	
1752		p = 0x00010002;
1753	
1754	     Please note that GCC really does use this sort of optimization,
1755	     which is not surprising given that it would likely take more
1756	     than two instructions to build the constant and then store it.
1757	     This optimization can therefore be a win in single-threaded code.
1758	     In fact, a recent bug (since fixed) caused GCC to incorrectly use
1759	     this optimization in a volatile store.  In the absence of such bugs,
1760	     use of WRITE_ONCE() prevents store tearing in the following example:
1761	
1762		WRITE_ONCE(p, 0x00010002);
1763	
1764	     Use of packed structures can also result in load and store tearing,
1765	     as in this example:
1766	
1767		struct __attribute__((__packed__)) foo {
1768			short a;
1769			int b;
1770			short c;
1771		};
1772		struct foo foo1, foo2;
1773		...
1774	
1775		foo2.a = foo1.a;
1776		foo2.b = foo1.b;
1777		foo2.c = foo1.c;
1778	
1779	     Because there are no READ_ONCE() or WRITE_ONCE() wrappers and no
1780	     volatile markings, the compiler would be well within its rights to
1781	     implement these three assignment statements as a pair of 32-bit
1782	     loads followed by a pair of 32-bit stores.  This would result in
1783	     load tearing on 'foo1.b' and store tearing on 'foo2.b'.  READ_ONCE()
1784	     and WRITE_ONCE() again prevent tearing in this example:
1785	
1786		foo2.a = foo1.a;
1787		WRITE_ONCE(foo2.b, READ_ONCE(foo1.b));
1788		foo2.c = foo1.c;
1789	
1790	All that aside, it is never necessary to use READ_ONCE() and
1791	WRITE_ONCE() on a variable that has been marked volatile.  For example,
1792	because 'jiffies' is marked volatile, it is never necessary to
1793	say READ_ONCE(jiffies).  The reason for this is that READ_ONCE() and
1794	WRITE_ONCE() are implemented as volatile casts, which has no effect when
1795	its argument is already marked volatile.
1796	
1797	Please note that these compiler barriers have no direct effect on the CPU,
1798	which may then reorder things however it wishes.
1799	
1800	
1801	CPU MEMORY BARRIERS
1802	-------------------
1803	
1804	The Linux kernel has eight basic CPU memory barriers:
1805	
1806		TYPE		MANDATORY		SMP CONDITIONAL
1807		===============	=======================	===========================
1808		GENERAL		mb()			smp_mb()
1809		WRITE		wmb()			smp_wmb()
1810		READ		rmb()			smp_rmb()
1811		DATA DEPENDENCY	read_barrier_depends()	smp_read_barrier_depends()
1812	
1813	
1814	All memory barriers except the data dependency barriers imply a compiler
1815	barrier.  Data dependencies do not impose any additional compiler ordering.
1816	
1817	Aside: In the case of data dependencies, the compiler would be expected
1818	to issue the loads in the correct order (eg. `a[b]` would have to load
1819	the value of b before loading a[b]), however there is no guarantee in
1820	the C specification that the compiler may not speculate the value of b
1821	(eg. is equal to 1) and load a before b (eg. tmp = a[1]; if (b != 1)
1822	tmp = a[b]; ).  There is also the problem of a compiler reloading b after
1823	having loaded a[b], thus having a newer copy of b than a[b].  A consensus
1824	has not yet been reached about these problems, however the READ_ONCE()
1825	macro is a good place to start looking.
1826	
1827	SMP memory barriers are reduced to compiler barriers on uniprocessor compiled
1828	systems because it is assumed that a CPU will appear to be self-consistent,
1829	and will order overlapping accesses correctly with respect to itself.
1830	However, see the subsection on "Virtual Machine Guests" below.
1831	
1832	[!] Note that SMP memory barriers _must_ be used to control the ordering of
1833	references to shared memory on SMP systems, though the use of locking instead
1834	is sufficient.
1835	
1836	Mandatory barriers should not be used to control SMP effects, since mandatory
1837	barriers impose unnecessary overhead on both SMP and UP systems. They may,
1838	however, be used to control MMIO effects on accesses through relaxed memory I/O
1839	windows.  These barriers are required even on non-SMP systems as they affect
1840	the order in which memory operations appear to a device by prohibiting both the
1841	compiler and the CPU from reordering them.
1842	
1843	
1844	There are some more advanced barrier functions:
1845	
1846	 (*) smp_store_mb(var, value)
1847	
1848	     This assigns the value to the variable and then inserts a full memory
1849	     barrier after it.  It isn't guaranteed to insert anything more than a
1850	     compiler barrier in a UP compilation.
1851	
1852	
1853	 (*) smp_mb__before_atomic();
1854	 (*) smp_mb__after_atomic();
1855	
1856	     These are for use with atomic (such as add, subtract, increment and
1857	     decrement) functions that don't return a value, especially when used for
1858	     reference counting.  These functions do not imply memory barriers.
1859	
1860	     These are also used for atomic bitop functions that do not return a
1861	     value (such as set_bit and clear_bit).
1862	
1863	     As an example, consider a piece of code that marks an object as being dead
1864	     and then decrements the object's reference count:
1865	
1866		obj->dead = 1;
1867		smp_mb__before_atomic();
1868		atomic_dec(&obj->ref_count);
1869	
1870	     This makes sure that the death mark on the object is perceived to be set
1871	     *before* the reference counter is decremented.
1872	
1873	     See Documentation/atomic_ops.txt for more information.  See the "Atomic
1874	     operations" subsection for information on where to use these.
1875	
1876	
1877	 (*) lockless_dereference();
1878	
1879	     This can be thought of as a pointer-fetch wrapper around the
1880	     smp_read_barrier_depends() data-dependency barrier.
1881	
1882	     This is also similar to rcu_dereference(), but in cases where
1883	     object lifetime is handled by some mechanism other than RCU, for
1884	     example, when the objects removed only when the system goes down.
1885	     In addition, lockless_dereference() is used in some data structures
1886	     that can be used both with and without RCU.
1887	
1888	
1889	 (*) dma_wmb();
1890	 (*) dma_rmb();
1891	
1892	     These are for use with consistent memory to guarantee the ordering
1893	     of writes or reads of shared memory accessible to both the CPU and a
1894	     DMA capable device.
1895	
1896	     For example, consider a device driver that shares memory with a device
1897	     and uses a descriptor status value to indicate if the descriptor belongs
1898	     to the device or the CPU, and a doorbell to notify it when new
1899	     descriptors are available:
1900	
1901		if (desc->status != DEVICE_OWN) {
1902			/* do not read data until we own descriptor */
1903			dma_rmb();
1904	
1905			/* read/modify data */
1906			read_data = desc->data;
1907			desc->data = write_data;
1908	
1909			/* flush modifications before status update */
1910			dma_wmb();
1911	
1912			/* assign ownership */
1913			desc->status = DEVICE_OWN;
1914	
1915			/* force memory to sync before notifying device via MMIO */
1916			wmb();
1917	
1918			/* notify device of new descriptors */
1919			writel(DESC_NOTIFY, doorbell);
1920		}
1921	
1922	     The dma_rmb() allows us guarantee the device has released ownership
1923	     before we read the data from the descriptor, and the dma_wmb() allows
1924	     us to guarantee the data is written to the descriptor before the device
1925	     can see it now has ownership.  The wmb() is needed to guarantee that the
1926	     cache coherent memory writes have completed before attempting a write to
1927	     the cache incoherent MMIO region.
1928	
1929	     See Documentation/DMA-API.txt for more information on consistent memory.
1930	
1931	MMIO WRITE BARRIER
1932	------------------
1933	
1934	The Linux kernel also has a special barrier for use with memory-mapped I/O
1935	writes:
1936	
1937		mmiowb();
1938	
1939	This is a variation on the mandatory write barrier that causes writes to weakly
1940	ordered I/O regions to be partially ordered.  Its effects may go beyond the
1941	CPU->Hardware interface and actually affect the hardware at some level.
1942	
1943	See the subsection "Acquires vs I/O accesses" for more information.
1944	
1945	
1946	===============================
1947	IMPLICIT KERNEL MEMORY BARRIERS
1948	===============================
1949	
1950	Some of the other functions in the linux kernel imply memory barriers, amongst
1951	which are locking and scheduling functions.
1952	
1953	This specification is a _minimum_ guarantee; any particular architecture may
1954	provide more substantial guarantees, but these may not be relied upon outside
1955	of arch specific code.
1956	
1957	
1958	LOCK ACQUISITION FUNCTIONS
1959	--------------------------
1960	
1961	The Linux kernel has a number of locking constructs:
1962	
1963	 (*) spin locks
1964	 (*) R/W spin locks
1965	 (*) mutexes
1966	 (*) semaphores
1967	 (*) R/W semaphores
1968	
1969	In all cases there are variants on "ACQUIRE" operations and "RELEASE" operations
1970	for each construct.  These operations all imply certain barriers:
1971	
1972	 (1) ACQUIRE operation implication:
1973	
1974	     Memory operations issued after the ACQUIRE will be completed after the
1975	     ACQUIRE operation has completed.
1976	
1977	     Memory operations issued before the ACQUIRE may be completed after
1978	     the ACQUIRE operation has completed.  An smp_mb__before_spinlock(),
1979	     combined with a following ACQUIRE, orders prior stores against
1980	     subsequent loads and stores.  Note that this is weaker than smp_mb()!
1981	     The smp_mb__before_spinlock() primitive is free on many architectures.
1982	
1983	 (2) RELEASE operation implication:
1984	
1985	     Memory operations issued before the RELEASE will be completed before the
1986	     RELEASE operation has completed.
1987	
1988	     Memory operations issued after the RELEASE may be completed before the
1989	     RELEASE operation has completed.
1990	
1991	 (3) ACQUIRE vs ACQUIRE implication:
1992	
1993	     All ACQUIRE operations issued before another ACQUIRE operation will be
1994	     completed before that ACQUIRE operation.
1995	
1996	 (4) ACQUIRE vs RELEASE implication:
1997	
1998	     All ACQUIRE operations issued before a RELEASE operation will be
1999	     completed before the RELEASE operation.
2000	
2001	 (5) Failed conditional ACQUIRE implication:
2002	
2003	     Certain locking variants of the ACQUIRE operation may fail, either due to
2004	     being unable to get the lock immediately, or due to receiving an unblocked
2005	     signal whilst asleep waiting for the lock to become available.  Failed
2006	     locks do not imply any sort of barrier.
2007	
2008	[!] Note: one of the consequences of lock ACQUIREs and RELEASEs being only
2009	one-way barriers is that the effects of instructions outside of a critical
2010	section may seep into the inside of the critical section.
2011	
2012	An ACQUIRE followed by a RELEASE may not be assumed to be full memory barrier
2013	because it is possible for an access preceding the ACQUIRE to happen after the
2014	ACQUIRE, and an access following the RELEASE to happen before the RELEASE, and
2015	the two accesses can themselves then cross:
2016	
2017		*A = a;
2018		ACQUIRE M
2019		RELEASE M
2020		*B = b;
2021	
2022	may occur as:
2023	
2024		ACQUIRE M, STORE *B, STORE *A, RELEASE M
2025	
2026	When the ACQUIRE and RELEASE are a lock acquisition and release,
2027	respectively, this same reordering can occur if the lock's ACQUIRE and
2028	RELEASE are to the same lock variable, but only from the perspective of
2029	another CPU not holding that lock.  In short, a ACQUIRE followed by an
2030	RELEASE may -not- be assumed to be a full memory barrier.
2031	
2032	Similarly, the reverse case of a RELEASE followed by an ACQUIRE does
2033	not imply a full memory barrier.  Therefore, the CPU's execution of the
2034	critical sections corresponding to the RELEASE and the ACQUIRE can cross,
2035	so that:
2036	
2037		*A = a;
2038		RELEASE M
2039		ACQUIRE N
2040		*B = b;
2041	
2042	could occur as:
2043	
2044		ACQUIRE N, STORE *B, STORE *A, RELEASE M
2045	
2046	It might appear that this reordering could introduce a deadlock.
2047	However, this cannot happen because if such a deadlock threatened,
2048	the RELEASE would simply complete, thereby avoiding the deadlock.
2049	
2050		Why does this work?
2051	
2052		One key point is that we are only talking about the CPU doing
2053		the reordering, not the compiler.  If the compiler (or, for
2054		that matter, the developer) switched the operations, deadlock
2055		-could- occur.
2056	
2057		But suppose the CPU reordered the operations.  In this case,
2058		the unlock precedes the lock in the assembly code.  The CPU
2059		simply elected to try executing the later lock operation first.
2060		If there is a deadlock, this lock operation will simply spin (or
2061		try to sleep, but more on that later).	The CPU will eventually
2062		execute the unlock operation (which preceded the lock operation
2063		in the assembly code), which will unravel the potential deadlock,
2064		allowing the lock operation to succeed.
2065	
2066		But what if the lock is a sleeplock?  In that case, the code will
2067		try to enter the scheduler, where it will eventually encounter
2068		a memory barrier, which will force the earlier unlock operation
2069		to complete, again unraveling the deadlock.  There might be
2070		a sleep-unlock race, but the locking primitive needs to resolve
2071		such races properly in any case.
2072	
2073	Locks and semaphores may not provide any guarantee of ordering on UP compiled
2074	systems, and so cannot be counted on in such a situation to actually achieve
2075	anything at all - especially with respect to I/O accesses - unless combined
2076	with interrupt disabling operations.
2077	
2078	See also the section on "Inter-CPU locking barrier effects".
2079	
2080	
2081	As an example, consider the following:
2082	
2083		*A = a;
2084		*B = b;
2085		ACQUIRE
2086		*C = c;
2087		*D = d;
2088		RELEASE
2089		*E = e;
2090		*F = f;
2091	
2092	The following sequence of events is acceptable:
2093	
2094		ACQUIRE, {*F,*A}, *E, {*C,*D}, *B, RELEASE
2095	
2096		[+] Note that {*F,*A} indicates a combined access.
2097	
2098	But none of the following are:
2099	
2100		{*F,*A}, *B,	ACQUIRE, *C, *D,	RELEASE, *E
2101		*A, *B, *C,	ACQUIRE, *D,		RELEASE, *E, *F
2102		*A, *B,		ACQUIRE, *C,		RELEASE, *D, *E, *F
2103		*B,		ACQUIRE, *C, *D,	RELEASE, {*F,*A}, *E
2104	
2105	
2106	
2107	INTERRUPT DISABLING FUNCTIONS
2108	-----------------------------
2109	
2110	Functions that disable interrupts (ACQUIRE equivalent) and enable interrupts
2111	(RELEASE equivalent) will act as compiler barriers only.  So if memory or I/O
2112	barriers are required in such a situation, they must be provided from some
2113	other means.
2114	
2115	
2116	SLEEP AND WAKE-UP FUNCTIONS
2117	---------------------------
2118	
2119	Sleeping and waking on an event flagged in global data can be viewed as an
2120	interaction between two pieces of data: the task state of the task waiting for
2121	the event and the global data used to indicate the event.  To make sure that
2122	these appear to happen in the right order, the primitives to begin the process
2123	of going to sleep, and the primitives to initiate a wake up imply certain
2124	barriers.
2125	
2126	Firstly, the sleeper normally follows something like this sequence of events:
2127	
2128		for (;;) {
2129			set_current_state(TASK_UNINTERRUPTIBLE);
2130			if (event_indicated)
2131				break;
2132			schedule();
2133		}
2134	
2135	A general memory barrier is interpolated automatically by set_current_state()
2136	after it has altered the task state:
2137	
2138		CPU 1
2139		===============================
2140		set_current_state();
2141		  smp_store_mb();
2142		    STORE current->state
2143		    <general barrier>
2144		LOAD event_indicated
2145	
2146	set_current_state() may be wrapped by:
2147	
2148		prepare_to_wait();
2149		prepare_to_wait_exclusive();
2150	
2151	which therefore also imply a general memory barrier after setting the state.
2152	The whole sequence above is available in various canned forms, all of which
2153	interpolate the memory barrier in the right place:
2154	
2155		wait_event();
2156		wait_event_interruptible();
2157		wait_event_interruptible_exclusive();
2158		wait_event_interruptible_timeout();
2159		wait_event_killable();
2160		wait_event_timeout();
2161		wait_on_bit();
2162		wait_on_bit_lock();
2163	
2164	
2165	Secondly, code that performs a wake up normally follows something like this:
2166	
2167		event_indicated = 1;
2168		wake_up(&event_wait_queue);
2169	
2170	or:
2171	
2172		event_indicated = 1;
2173		wake_up_process(event_daemon);
2174	
2175	A write memory barrier is implied by wake_up() and co.  if and only if they
2176	wake something up.  The barrier occurs before the task state is cleared, and so
2177	sits between the STORE to indicate the event and the STORE to set TASK_RUNNING:
2178	
2179		CPU 1				CPU 2
2180		===============================	===============================
2181		set_current_state();		STORE event_indicated
2182		  smp_store_mb();		wake_up();
2183		    STORE current->state	  <write barrier>
2184		    <general barrier>		  STORE current->state
2185		LOAD event_indicated
2186	
2187	To repeat, this write memory barrier is present if and only if something
2188	is actually awakened.  To see this, consider the following sequence of
2189	events, where X and Y are both initially zero:
2190	
2191		CPU 1				CPU 2
2192		===============================	===============================
2193		X = 1;				STORE event_indicated
2194		smp_mb();			wake_up();
2195		Y = 1;				wait_event(wq, Y == 1);
2196		wake_up();			  load from Y sees 1, no memory barrier
2197						load from X might see 0
2198	
2199	In contrast, if a wakeup does occur, CPU 2's load from X would be guaranteed
2200	to see 1.
2201	
2202	The available waker functions include:
2203	
2204		complete();
2205		wake_up();
2206		wake_up_all();
2207		wake_up_bit();
2208		wake_up_interruptible();
2209		wake_up_interruptible_all();
2210		wake_up_interruptible_nr();
2211		wake_up_interruptible_poll();
2212		wake_up_interruptible_sync();
2213		wake_up_interruptible_sync_poll();
2214		wake_up_locked();
2215		wake_up_locked_poll();
2216		wake_up_nr();
2217		wake_up_poll();
2218		wake_up_process();
2219	
2220	
2221	[!] Note that the memory barriers implied by the sleeper and the waker do _not_
2222	order multiple stores before the wake-up with respect to loads of those stored
2223	values after the sleeper has called set_current_state().  For instance, if the
2224	sleeper does:
2225	
2226		set_current_state(TASK_INTERRUPTIBLE);
2227		if (event_indicated)
2228			break;
2229		__set_current_state(TASK_RUNNING);
2230		do_something(my_data);
2231	
2232	and the waker does:
2233	
2234		my_data = value;
2235		event_indicated = 1;
2236		wake_up(&event_wait_queue);
2237	
2238	there's no guarantee that the change to event_indicated will be perceived by
2239	the sleeper as coming after the change to my_data.  In such a circumstance, the
2240	code on both sides must interpolate its own memory barriers between the
2241	separate data accesses.  Thus the above sleeper ought to do:
2242	
2243		set_current_state(TASK_INTERRUPTIBLE);
2244		if (event_indicated) {
2245			smp_rmb();
2246			do_something(my_data);
2247		}
2248	
2249	and the waker should do:
2250	
2251		my_data = value;
2252		smp_wmb();
2253		event_indicated = 1;
2254		wake_up(&event_wait_queue);
2255	
2256	
2257	MISCELLANEOUS FUNCTIONS
2258	-----------------------
2259	
2260	Other functions that imply barriers:
2261	
2262	 (*) schedule() and similar imply full memory barriers.
2263	
2264	
2265	===================================
2266	INTER-CPU ACQUIRING BARRIER EFFECTS
2267	===================================
2268	
2269	On SMP systems locking primitives give a more substantial form of barrier: one
2270	that does affect memory access ordering on other CPUs, within the context of
2271	conflict on any particular lock.
2272	
2273	
2274	ACQUIRES VS MEMORY ACCESSES
2275	---------------------------
2276	
2277	Consider the following: the system has a pair of spinlocks (M) and (Q), and
2278	three CPUs; then should the following sequence of events occur:
2279	
2280		CPU 1				CPU 2
2281		===============================	===============================
2282		WRITE_ONCE(*A, a);		WRITE_ONCE(*E, e);
2283		ACQUIRE M			ACQUIRE Q
2284		WRITE_ONCE(*B, b);		WRITE_ONCE(*F, f);
2285		WRITE_ONCE(*C, c);		WRITE_ONCE(*G, g);
2286		RELEASE M			RELEASE Q
2287		WRITE_ONCE(*D, d);		WRITE_ONCE(*H, h);
2288	
2289	Then there is no guarantee as to what order CPU 3 will see the accesses to *A
2290	through *H occur in, other than the constraints imposed by the separate locks
2291	on the separate CPUs.  It might, for example, see:
2292	
2293		*E, ACQUIRE M, ACQUIRE Q, *G, *C, *F, *A, *B, RELEASE Q, *D, *H, RELEASE M
2294	
2295	But it won't see any of:
2296	
2297		*B, *C or *D preceding ACQUIRE M
2298		*A, *B or *C following RELEASE M
2299		*F, *G or *H preceding ACQUIRE Q
2300		*E, *F or *G following RELEASE Q
2301	
2302	
2303	
2304	ACQUIRES VS I/O ACCESSES
2305	------------------------
2306	
2307	Under certain circumstances (especially involving NUMA), I/O accesses within
2308	two spinlocked sections on two different CPUs may be seen as interleaved by the
2309	PCI bridge, because the PCI bridge does not necessarily participate in the
2310	cache-coherence protocol, and is therefore incapable of issuing the required
2311	read memory barriers.
2312	
2313	For example:
2314	
2315		CPU 1				CPU 2
2316		===============================	===============================
2317		spin_lock(Q)
2318		writel(0, ADDR)
2319		writel(1, DATA);
2320		spin_unlock(Q);
2321						spin_lock(Q);
2322						writel(4, ADDR);
2323						writel(5, DATA);
2324						spin_unlock(Q);
2325	
2326	may be seen by the PCI bridge as follows:
2327	
2328		STORE *ADDR = 0, STORE *ADDR = 4, STORE *DATA = 1, STORE *DATA = 5
2329	
2330	which would probably cause the hardware to malfunction.
2331	
2332	
2333	What is necessary here is to intervene with an mmiowb() before dropping the
2334	spinlock, for example:
2335	
2336		CPU 1				CPU 2
2337		===============================	===============================
2338		spin_lock(Q)
2339		writel(0, ADDR)
2340		writel(1, DATA);
2341		mmiowb();
2342		spin_unlock(Q);
2343						spin_lock(Q);
2344						writel(4, ADDR);
2345						writel(5, DATA);
2346						mmiowb();
2347						spin_unlock(Q);
2348	
2349	this will ensure that the two stores issued on CPU 1 appear at the PCI bridge
2350	before either of the stores issued on CPU 2.
2351	
2352	
2353	Furthermore, following a store by a load from the same device obviates the need
2354	for the mmiowb(), because the load forces the store to complete before the load
2355	is performed:
2356	
2357		CPU 1				CPU 2
2358		===============================	===============================
2359		spin_lock(Q)
2360		writel(0, ADDR)
2361		a = readl(DATA);
2362		spin_unlock(Q);
2363						spin_lock(Q);
2364						writel(4, ADDR);
2365						b = readl(DATA);
2366						spin_unlock(Q);
2367	
2368	
2369	See Documentation/DocBook/deviceiobook.tmpl for more information.
2370	
2371	
2372	=================================
2373	WHERE ARE MEMORY BARRIERS NEEDED?
2374	=================================
2375	
2376	Under normal operation, memory operation reordering is generally not going to
2377	be a problem as a single-threaded linear piece of code will still appear to
2378	work correctly, even if it's in an SMP kernel.  There are, however, four
2379	circumstances in which reordering definitely _could_ be a problem:
2380	
2381	 (*) Interprocessor interaction.
2382	
2383	 (*) Atomic operations.
2384	
2385	 (*) Accessing devices.
2386	
2387	 (*) Interrupts.
2388	
2389	
2390	INTERPROCESSOR INTERACTION
2391	--------------------------
2392	
2393	When there's a system with more than one processor, more than one CPU in the
2394	system may be working on the same data set at the same time.  This can cause
2395	synchronisation problems, and the usual way of dealing with them is to use
2396	locks.  Locks, however, are quite expensive, and so it may be preferable to
2397	operate without the use of a lock if at all possible.  In such a case
2398	operations that affect both CPUs may have to be carefully ordered to prevent
2399	a malfunction.
2400	
2401	Consider, for example, the R/W semaphore slow path.  Here a waiting process is
2402	queued on the semaphore, by virtue of it having a piece of its stack linked to
2403	the semaphore's list of waiting processes:
2404	
2405		struct rw_semaphore {
2406			...
2407			spinlock_t lock;
2408			struct list_head waiters;
2409		};
2410	
2411		struct rwsem_waiter {
2412			struct list_head list;
2413			struct task_struct *task;
2414		};
2415	
2416	To wake up a particular waiter, the up_read() or up_write() functions have to:
2417	
2418	 (1) read the next pointer from this waiter's record to know as to where the
2419	     next waiter record is;
2420	
2421	 (2) read the pointer to the waiter's task structure;
2422	
2423	 (3) clear the task pointer to tell the waiter it has been given the semaphore;
2424	
2425	 (4) call wake_up_process() on the task; and
2426	
2427	 (5) release the reference held on the waiter's task struct.
2428	
2429	In other words, it has to perform this sequence of events:
2430	
2431		LOAD waiter->list.next;
2432		LOAD waiter->task;
2433		STORE waiter->task;
2434		CALL wakeup
2435		RELEASE task
2436	
2437	and if any of these steps occur out of order, then the whole thing may
2438	malfunction.
2439	
2440	Once it has queued itself and dropped the semaphore lock, the waiter does not
2441	get the lock again; it instead just waits for its task pointer to be cleared
2442	before proceeding.  Since the record is on the waiter's stack, this means that
2443	if the task pointer is cleared _before_ the next pointer in the list is read,
2444	another CPU might start processing the waiter and might clobber the waiter's
2445	stack before the up*() function has a chance to read the next pointer.
2446	
2447	Consider then what might happen to the above sequence of events:
2448	
2449		CPU 1				CPU 2
2450		===============================	===============================
2451						down_xxx()
2452						Queue waiter
2453						Sleep
2454		up_yyy()
2455		LOAD waiter->task;
2456		STORE waiter->task;
2457						Woken up by other event
2458		<preempt>
2459						Resume processing
2460						down_xxx() returns
2461						call foo()
2462						foo() clobbers *waiter
2463		</preempt>
2464		LOAD waiter->list.next;
2465		--- OOPS ---
2466	
2467	This could be dealt with using the semaphore lock, but then the down_xxx()
2468	function has to needlessly get the spinlock again after being woken up.
2469	
2470	The way to deal with this is to insert a general SMP memory barrier:
2471	
2472		LOAD waiter->list.next;
2473		LOAD waiter->task;
2474		smp_mb();
2475		STORE waiter->task;
2476		CALL wakeup
2477		RELEASE task
2478	
2479	In this case, the barrier makes a guarantee that all memory accesses before the
2480	barrier will appear to happen before all the memory accesses after the barrier
2481	with respect to the other CPUs on the system.  It does _not_ guarantee that all
2482	the memory accesses before the barrier will be complete by the time the barrier
2483	instruction itself is complete.
2484	
2485	On a UP system - where this wouldn't be a problem - the smp_mb() is just a
2486	compiler barrier, thus making sure the compiler emits the instructions in the
2487	right order without actually intervening in the CPU.  Since there's only one
2488	CPU, that CPU's dependency ordering logic will take care of everything else.
2489	
2490	
2491	ATOMIC OPERATIONS
2492	-----------------
2493	
2494	Whilst they are technically interprocessor interaction considerations, atomic
2495	operations are noted specially as some of them imply full memory barriers and
2496	some don't, but they're very heavily relied on as a group throughout the
2497	kernel.
2498	
2499	Any atomic operation that modifies some state in memory and returns information
2500	about the state (old or new) implies an SMP-conditional general memory barrier
2501	(smp_mb()) on each side of the actual operation (with the exception of
2502	explicit lock operations, described later).  These include:
2503	
2504		xchg();
2505		atomic_xchg();			atomic_long_xchg();
2506		atomic_inc_return();		atomic_long_inc_return();
2507		atomic_dec_return();		atomic_long_dec_return();
2508		atomic_add_return();		atomic_long_add_return();
2509		atomic_sub_return();		atomic_long_sub_return();
2510		atomic_inc_and_test();		atomic_long_inc_and_test();
2511		atomic_dec_and_test();		atomic_long_dec_and_test();
2512		atomic_sub_and_test();		atomic_long_sub_and_test();
2513		atomic_add_negative();		atomic_long_add_negative();
2514		test_and_set_bit();
2515		test_and_clear_bit();
2516		test_and_change_bit();
2517	
2518		/* when succeeds */
2519		cmpxchg();
2520		atomic_cmpxchg();		atomic_long_cmpxchg();
2521		atomic_add_unless();		atomic_long_add_unless();
2522	
2523	These are used for such things as implementing ACQUIRE-class and RELEASE-class
2524	operations and adjusting reference counters towards object destruction, and as
2525	such the implicit memory barrier effects are necessary.
2526	
2527	
2528	The following operations are potential problems as they do _not_ imply memory
2529	barriers, but might be used for implementing such things as RELEASE-class
2530	operations:
2531	
2532		atomic_set();
2533		set_bit();
2534		clear_bit();
2535		change_bit();
2536	
2537	With these the appropriate explicit memory barrier should be used if necessary
2538	(smp_mb__before_atomic() for instance).
2539	
2540	
2541	The following also do _not_ imply memory barriers, and so may require explicit
2542	memory barriers under some circumstances (smp_mb__before_atomic() for
2543	instance):
2544	
2545		atomic_add();
2546		atomic_sub();
2547		atomic_inc();
2548		atomic_dec();
2549	
2550	If they're used for statistics generation, then they probably don't need memory
2551	barriers, unless there's a coupling between statistical data.
2552	
2553	If they're used for reference counting on an object to control its lifetime,
2554	they probably don't need memory barriers because either the reference count
2555	will be adjusted inside a locked section, or the caller will already hold
2556	sufficient references to make the lock, and thus a memory barrier unnecessary.
2557	
2558	If they're used for constructing a lock of some description, then they probably
2559	do need memory barriers as a lock primitive generally has to do things in a
2560	specific order.
2561	
2562	Basically, each usage case has to be carefully considered as to whether memory
2563	barriers are needed or not.
2564	
2565	The following operations are special locking primitives:
2566	
2567		test_and_set_bit_lock();
2568		clear_bit_unlock();
2569		__clear_bit_unlock();
2570	
2571	These implement ACQUIRE-class and RELEASE-class operations.  These should be
2572	used in preference to other operations when implementing locking primitives,
2573	because their implementations can be optimised on many architectures.
2574	
2575	[!] Note that special memory barrier primitives are available for these
2576	situations because on some CPUs the atomic instructions used imply full memory
2577	barriers, and so barrier instructions are superfluous in conjunction with them,
2578	and in such cases the special barrier primitives will be no-ops.
2579	
2580	See Documentation/atomic_ops.txt for more information.
2581	
2582	
2583	ACCESSING DEVICES
2584	-----------------
2585	
2586	Many devices can be memory mapped, and so appear to the CPU as if they're just
2587	a set of memory locations.  To control such a device, the driver usually has to
2588	make the right memory accesses in exactly the right order.
2589	
2590	However, having a clever CPU or a clever compiler creates a potential problem
2591	in that the carefully sequenced accesses in the driver code won't reach the
2592	device in the requisite order if the CPU or the compiler thinks it is more
2593	efficient to reorder, combine or merge accesses - something that would cause
2594	the device to malfunction.
2595	
2596	Inside of the Linux kernel, I/O should be done through the appropriate accessor
2597	routines - such as inb() or writel() - which know how to make such accesses
2598	appropriately sequential.  Whilst this, for the most part, renders the explicit
2599	use of memory barriers unnecessary, there are a couple of situations where they
2600	might be needed:
2601	
2602	 (1) On some systems, I/O stores are not strongly ordered across all CPUs, and
2603	     so for _all_ general drivers locks should be used and mmiowb() must be
2604	     issued prior to unlocking the critical section.
2605	
2606	 (2) If the accessor functions are used to refer to an I/O memory window with
2607	     relaxed memory access properties, then _mandatory_ memory barriers are
2608	     required to enforce ordering.
2609	
2610	See Documentation/DocBook/deviceiobook.tmpl for more information.
2611	
2612	
2613	INTERRUPTS
2614	----------
2615	
2616	A driver may be interrupted by its own interrupt service routine, and thus the
2617	two parts of the driver may interfere with each other's attempts to control or
2618	access the device.
2619	
2620	This may be alleviated - at least in part - by disabling local interrupts (a
2621	form of locking), such that the critical operations are all contained within
2622	the interrupt-disabled section in the driver.  Whilst the driver's interrupt
2623	routine is executing, the driver's core may not run on the same CPU, and its
2624	interrupt is not permitted to happen again until the current interrupt has been
2625	handled, thus the interrupt handler does not need to lock against that.
2626	
2627	However, consider a driver that was talking to an ethernet card that sports an
2628	address register and a data register.  If that driver's core talks to the card
2629	under interrupt-disablement and then the driver's interrupt handler is invoked:
2630	
2631		LOCAL IRQ DISABLE
2632		writew(ADDR, 3);
2633		writew(DATA, y);
2634		LOCAL IRQ ENABLE
2635		<interrupt>
2636		writew(ADDR, 4);
2637		q = readw(DATA);
2638		</interrupt>
2639	
2640	The store to the data register might happen after the second store to the
2641	address register if ordering rules are sufficiently relaxed:
2642	
2643		STORE *ADDR = 3, STORE *ADDR = 4, STORE *DATA = y, q = LOAD *DATA
2644	
2645	
2646	If ordering rules are relaxed, it must be assumed that accesses done inside an
2647	interrupt disabled section may leak outside of it and may interleave with
2648	accesses performed in an interrupt - and vice versa - unless implicit or
2649	explicit barriers are used.
2650	
2651	Normally this won't be a problem because the I/O accesses done inside such
2652	sections will include synchronous load operations on strictly ordered I/O
2653	registers that form implicit I/O barriers.  If this isn't sufficient then an
2654	mmiowb() may need to be used explicitly.
2655	
2656	
2657	A similar situation may occur between an interrupt routine and two routines
2658	running on separate CPUs that communicate with each other.  If such a case is
2659	likely, then interrupt-disabling locks should be used to guarantee ordering.
2660	
2661	
2662	==========================
2663	KERNEL I/O BARRIER EFFECTS
2664	==========================
2665	
2666	When accessing I/O memory, drivers should use the appropriate accessor
2667	functions:
2668	
2669	 (*) inX(), outX():
2670	
2671	     These are intended to talk to I/O space rather than memory space, but
2672	     that's primarily a CPU-specific concept.  The i386 and x86_64 processors
2673	     do indeed have special I/O space access cycles and instructions, but many
2674	     CPUs don't have such a concept.
2675	
2676	     The PCI bus, amongst others, defines an I/O space concept which - on such
2677	     CPUs as i386 and x86_64 - readily maps to the CPU's concept of I/O
2678	     space.  However, it may also be mapped as a virtual I/O space in the CPU's
2679	     memory map, particularly on those CPUs that don't support alternate I/O
2680	     spaces.
2681	
2682	     Accesses to this space may be fully synchronous (as on i386), but
2683	     intermediary bridges (such as the PCI host bridge) may not fully honour
2684	     that.
2685	
2686	     They are guaranteed to be fully ordered with respect to each other.
2687	
2688	     They are not guaranteed to be fully ordered with respect to other types of
2689	     memory and I/O operation.
2690	
2691	 (*) readX(), writeX():
2692	
2693	     Whether these are guaranteed to be fully ordered and uncombined with
2694	     respect to each other on the issuing CPU depends on the characteristics
2695	     defined for the memory window through which they're accessing.  On later
2696	     i386 architecture machines, for example, this is controlled by way of the
2697	     MTRR registers.
2698	
2699	     Ordinarily, these will be guaranteed to be fully ordered and uncombined,
2700	     provided they're not accessing a prefetchable device.
2701	
2702	     However, intermediary hardware (such as a PCI bridge) may indulge in
2703	     deferral if it so wishes; to flush a store, a load from the same location
2704	     is preferred[*], but a load from the same device or from configuration
2705	     space should suffice for PCI.
2706	
2707	     [*] NOTE! attempting to load from the same location as was written to may
2708		 cause a malfunction - consider the 16550 Rx/Tx serial registers for
2709		 example.
2710	
2711	     Used with prefetchable I/O memory, an mmiowb() barrier may be required to
2712	     force stores to be ordered.
2713	
2714	     Please refer to the PCI specification for more information on interactions
2715	     between PCI transactions.
2716	
2717	 (*) readX_relaxed(), writeX_relaxed()
2718	
2719	     These are similar to readX() and writeX(), but provide weaker memory
2720	     ordering guarantees.  Specifically, they do not guarantee ordering with
2721	     respect to normal memory accesses (e.g. DMA buffers) nor do they guarantee
2722	     ordering with respect to LOCK or UNLOCK operations.  If the latter is
2723	     required, an mmiowb() barrier can be used.  Note that relaxed accesses to
2724	     the same peripheral are guaranteed to be ordered with respect to each
2725	     other.
2726	
2727	 (*) ioreadX(), iowriteX()
2728	
2729	     These will perform appropriately for the type of access they're actually
2730	     doing, be it inX()/outX() or readX()/writeX().
2731	
2732	
2733	========================================
2734	ASSUMED MINIMUM EXECUTION ORDERING MODEL
2735	========================================
2736	
2737	It has to be assumed that the conceptual CPU is weakly-ordered but that it will
2738	maintain the appearance of program causality with respect to itself.  Some CPUs
2739	(such as i386 or x86_64) are more constrained than others (such as powerpc or
2740	frv), and so the most relaxed case (namely DEC Alpha) must be assumed outside
2741	of arch-specific code.
2742	
2743	This means that it must be considered that the CPU will execute its instruction
2744	stream in any order it feels like - or even in parallel - provided that if an
2745	instruction in the stream depends on an earlier instruction, then that
2746	earlier instruction must be sufficiently complete[*] before the later
2747	instruction may proceed; in other words: provided that the appearance of
2748	causality is maintained.
2749	
2750	 [*] Some instructions have more than one effect - such as changing the
2751	     condition codes, changing registers or changing memory - and different
2752	     instructions may depend on different effects.
2753	
2754	A CPU may also discard any instruction sequence that winds up having no
2755	ultimate effect.  For example, if two adjacent instructions both load an
2756	immediate value into the same register, the first may be discarded.
2757	
2758	
2759	Similarly, it has to be assumed that compiler might reorder the instruction
2760	stream in any way it sees fit, again provided the appearance of causality is
2761	maintained.
2762	
2763	
2764	============================
2765	THE EFFECTS OF THE CPU CACHE
2766	============================
2767	
2768	The way cached memory operations are perceived across the system is affected to
2769	a certain extent by the caches that lie between CPUs and memory, and by the
2770	memory coherence system that maintains the consistency of state in the system.
2771	
2772	As far as the way a CPU interacts with another part of the system through the
2773	caches goes, the memory system has to include the CPU's caches, and memory
2774	barriers for the most part act at the interface between the CPU and its cache
2775	(memory barriers logically act on the dotted line in the following diagram):
2776	
2777		    <--- CPU --->         :       <----------- Memory ----------->
2778		                          :
2779		+--------+    +--------+  :   +--------+    +-----------+
2780		|        |    |        |  :   |        |    |           |    +--------+
2781		|  CPU   |    | Memory |  :   | CPU    |    |           |    |        |
2782		|  Core  |--->| Access |----->| Cache  |<-->|           |    |        |
2783		|        |    | Queue  |  :   |        |    |           |--->| Memory |
2784		|        |    |        |  :   |        |    |           |    |        |
2785		+--------+    +--------+  :   +--------+    |           |    |        |
2786		                          :                 | Cache     |    +--------+
2787		                          :                 | Coherency |
2788		                          :                 | Mechanism |    +--------+
2789		+--------+    +--------+  :   +--------+    |           |    |	      |
2790		|        |    |        |  :   |        |    |           |    |        |
2791		|  CPU   |    | Memory |  :   | CPU    |    |           |--->| Device |
2792		|  Core  |--->| Access |----->| Cache  |<-->|           |    |        |
2793		|        |    | Queue  |  :   |        |    |           |    |        |
2794		|        |    |        |  :   |        |    |           |    +--------+
2795		+--------+    +--------+  :   +--------+    +-----------+
2796		                          :
2797		                          :
2798	
2799	Although any particular load or store may not actually appear outside of the
2800	CPU that issued it since it may have been satisfied within the CPU's own cache,
2801	it will still appear as if the full memory access had taken place as far as the
2802	other CPUs are concerned since the cache coherency mechanisms will migrate the
2803	cacheline over to the accessing CPU and propagate the effects upon conflict.
2804	
2805	The CPU core may execute instructions in any order it deems fit, provided the
2806	expected program causality appears to be maintained.  Some of the instructions
2807	generate load and store operations which then go into the queue of memory
2808	accesses to be performed.  The core may place these in the queue in any order
2809	it wishes, and continue execution until it is forced to wait for an instruction
2810	to complete.
2811	
2812	What memory barriers are concerned with is controlling the order in which
2813	accesses cross from the CPU side of things to the memory side of things, and
2814	the order in which the effects are perceived to happen by the other observers
2815	in the system.
2816	
2817	[!] Memory barriers are _not_ needed within a given CPU, as CPUs always see
2818	their own loads and stores as if they had happened in program order.
2819	
2820	[!] MMIO or other device accesses may bypass the cache system.  This depends on
2821	the properties of the memory window through which devices are accessed and/or
2822	the use of any special device communication instructions the CPU may have.
2823	
2824	
2825	CACHE COHERENCY
2826	---------------
2827	
2828	Life isn't quite as simple as it may appear above, however: for while the
2829	caches are expected to be coherent, there's no guarantee that that coherency
2830	will be ordered.  This means that whilst changes made on one CPU will
2831	eventually become visible on all CPUs, there's no guarantee that they will
2832	become apparent in the same order on those other CPUs.
2833	
2834	
2835	Consider dealing with a system that has a pair of CPUs (1 & 2), each of which
2836	has a pair of parallel data caches (CPU 1 has A/B, and CPU 2 has C/D):
2837	
2838		            :
2839		            :                          +--------+
2840		            :      +---------+         |        |
2841		+--------+  : +--->| Cache A |<------->|        |
2842		|        |  : |    +---------+         |        |
2843		|  CPU 1 |<---+                        |        |
2844		|        |  : |    +---------+         |        |
2845		+--------+  : +--->| Cache B |<------->|        |
2846		            :      +---------+         |        |
2847		            :                          | Memory |
2848		            :      +---------+         | System |
2849		+--------+  : +--->| Cache C |<------->|        |
2850		|        |  : |    +---------+         |        |
2851		|  CPU 2 |<---+                        |        |
2852		|        |  : |    +---------+         |        |
2853		+--------+  : +--->| Cache D |<------->|        |
2854		            :      +---------+         |        |
2855		            :                          +--------+
2856		            :
2857	
2858	Imagine the system has the following properties:
2859	
2860	 (*) an odd-numbered cache line may be in cache A, cache C or it may still be
2861	     resident in memory;
2862	
2863	 (*) an even-numbered cache line may be in cache B, cache D or it may still be
2864	     resident in memory;
2865	
2866	 (*) whilst the CPU core is interrogating one cache, the other cache may be
2867	     making use of the bus to access the rest of the system - perhaps to
2868	     displace a dirty cacheline or to do a speculative load;
2869	
2870	 (*) each cache has a queue of operations that need to be applied to that cache
2871	     to maintain coherency with the rest of the system;
2872	
2873	 (*) the coherency queue is not flushed by normal loads to lines already
2874	     present in the cache, even though the contents of the queue may
2875	     potentially affect those loads.
2876	
2877	Imagine, then, that two writes are made on the first CPU, with a write barrier
2878	between them to guarantee that they will appear to reach that CPU's caches in
2879	the requisite order:
2880	
2881		CPU 1		CPU 2		COMMENT
2882		===============	===============	=======================================
2883						u == 0, v == 1 and p == &u, q == &u
2884		v = 2;
2885		smp_wmb();			Make sure change to v is visible before
2886						 change to p
2887		<A:modify v=2>			v is now in cache A exclusively
2888		p = &v;
2889		<B:modify p=&v>			p is now in cache B exclusively
2890	
2891	The write memory barrier forces the other CPUs in the system to perceive that
2892	the local CPU's caches have apparently been updated in the correct order.  But
2893	now imagine that the second CPU wants to read those values:
2894	
2895		CPU 1		CPU 2		COMMENT
2896		===============	===============	=======================================
2897		...
2898				q = p;
2899				x = *q;
2900	
2901	The above pair of reads may then fail to happen in the expected order, as the
2902	cacheline holding p may get updated in one of the second CPU's caches whilst
2903	the update to the cacheline holding v is delayed in the other of the second
2904	CPU's caches by some other cache event:
2905	
2906		CPU 1		CPU 2		COMMENT
2907		===============	===============	=======================================
2908						u == 0, v == 1 and p == &u, q == &u
2909		v = 2;
2910		smp_wmb();
2911		<A:modify v=2>	<C:busy>
2912				<C:queue v=2>
2913		p = &v;		q = p;
2914				<D:request p>
2915		<B:modify p=&v>	<D:commit p=&v>
2916				<D:read p>
2917				x = *q;
2918				<C:read *q>	Reads from v before v updated in cache
2919				<C:unbusy>
2920				<C:commit v=2>
2921	
2922	Basically, whilst both cachelines will be updated on CPU 2 eventually, there's
2923	no guarantee that, without intervention, the order of update will be the same
2924	as that committed on CPU 1.
2925	
2926	
2927	To intervene, we need to interpolate a data dependency barrier or a read
2928	barrier between the loads.  This will force the cache to commit its coherency
2929	queue before processing any further requests:
2930	
2931		CPU 1		CPU 2		COMMENT
2932		===============	===============	=======================================
2933						u == 0, v == 1 and p == &u, q == &u
2934		v = 2;
2935		smp_wmb();
2936		<A:modify v=2>	<C:busy>
2937				<C:queue v=2>
2938		p = &v;		q = p;
2939				<D:request p>
2940		<B:modify p=&v>	<D:commit p=&v>
2941				<D:read p>
2942				smp_read_barrier_depends()
2943				<C:unbusy>
2944				<C:commit v=2>
2945				x = *q;
2946				<C:read *q>	Reads from v after v updated in cache
2947	
2948	
2949	This sort of problem can be encountered on DEC Alpha processors as they have a
2950	split cache that improves performance by making better use of the data bus.
2951	Whilst most CPUs do imply a data dependency barrier on the read when a memory
2952	access depends on a read, not all do, so it may not be relied on.
2953	
2954	Other CPUs may also have split caches, but must coordinate between the various
2955	cachelets for normal memory accesses.  The semantics of the Alpha removes the
2956	need for coordination in the absence of memory barriers.
2957	
2958	
2959	CACHE COHERENCY VS DMA
2960	----------------------
2961	
2962	Not all systems maintain cache coherency with respect to devices doing DMA.  In
2963	such cases, a device attempting DMA may obtain stale data from RAM because
2964	dirty cache lines may be resident in the caches of various CPUs, and may not
2965	have been written back to RAM yet.  To deal with this, the appropriate part of
2966	the kernel must flush the overlapping bits of cache on each CPU (and maybe
2967	invalidate them as well).
2968	
2969	In addition, the data DMA'd to RAM by a device may be overwritten by dirty
2970	cache lines being written back to RAM from a CPU's cache after the device has
2971	installed its own data, or cache lines present in the CPU's cache may simply
2972	obscure the fact that RAM has been updated, until at such time as the cacheline
2973	is discarded from the CPU's cache and reloaded.  To deal with this, the
2974	appropriate part of the kernel must invalidate the overlapping bits of the
2975	cache on each CPU.
2976	
2977	See Documentation/cachetlb.txt for more information on cache management.
2978	
2979	
2980	CACHE COHERENCY VS MMIO
2981	-----------------------
2982	
2983	Memory mapped I/O usually takes place through memory locations that are part of
2984	a window in the CPU's memory space that has different properties assigned than
2985	the usual RAM directed window.
2986	
2987	Amongst these properties is usually the fact that such accesses bypass the
2988	caching entirely and go directly to the device buses.  This means MMIO accesses
2989	may, in effect, overtake accesses to cached memory that were emitted earlier.
2990	A memory barrier isn't sufficient in such a case, but rather the cache must be
2991	flushed between the cached memory write and the MMIO access if the two are in
2992	any way dependent.
2993	
2994	
2995	=========================
2996	THE THINGS CPUS GET UP TO
2997	=========================
2998	
2999	A programmer might take it for granted that the CPU will perform memory
3000	operations in exactly the order specified, so that if the CPU is, for example,
3001	given the following piece of code to execute:
3002	
3003		a = READ_ONCE(*A);
3004		WRITE_ONCE(*B, b);
3005		c = READ_ONCE(*C);
3006		d = READ_ONCE(*D);
3007		WRITE_ONCE(*E, e);
3008	
3009	they would then expect that the CPU will complete the memory operation for each
3010	instruction before moving on to the next one, leading to a definite sequence of
3011	operations as seen by external observers in the system:
3012	
3013		LOAD *A, STORE *B, LOAD *C, LOAD *D, STORE *E.
3014	
3015	
3016	Reality is, of course, much messier.  With many CPUs and compilers, the above
3017	assumption doesn't hold because:
3018	
3019	 (*) loads are more likely to need to be completed immediately to permit
3020	     execution progress, whereas stores can often be deferred without a
3021	     problem;
3022	
3023	 (*) loads may be done speculatively, and the result discarded should it prove
3024	     to have been unnecessary;
3025	
3026	 (*) loads may be done speculatively, leading to the result having been fetched
3027	     at the wrong time in the expected sequence of events;
3028	
3029	 (*) the order of the memory accesses may be rearranged to promote better use
3030	     of the CPU buses and caches;
3031	
3032	 (*) loads and stores may be combined to improve performance when talking to
3033	     memory or I/O hardware that can do batched accesses of adjacent locations,
3034	     thus cutting down on transaction setup costs (memory and PCI devices may
3035	     both be able to do this); and
3036	
3037	 (*) the CPU's data cache may affect the ordering, and whilst cache-coherency
3038	     mechanisms may alleviate this - once the store has actually hit the cache
3039	     - there's no guarantee that the coherency management will be propagated in
3040	     order to other CPUs.
3041	
3042	So what another CPU, say, might actually observe from the above piece of code
3043	is:
3044	
3045		LOAD *A, ..., LOAD {*C,*D}, STORE *E, STORE *B
3046	
3047		(Where "LOAD {*C,*D}" is a combined load)
3048	
3049	
3050	However, it is guaranteed that a CPU will be self-consistent: it will see its
3051	_own_ accesses appear to be correctly ordered, without the need for a memory
3052	barrier.  For instance with the following code:
3053	
3054		U = READ_ONCE(*A);
3055		WRITE_ONCE(*A, V);
3056		WRITE_ONCE(*A, W);
3057		X = READ_ONCE(*A);
3058		WRITE_ONCE(*A, Y);
3059		Z = READ_ONCE(*A);
3060	
3061	and assuming no intervention by an external influence, it can be assumed that
3062	the final result will appear to be:
3063	
3064		U == the original value of *A
3065		X == W
3066		Z == Y
3067		*A == Y
3068	
3069	The code above may cause the CPU to generate the full sequence of memory
3070	accesses:
3071	
3072		U=LOAD *A, STORE *A=V, STORE *A=W, X=LOAD *A, STORE *A=Y, Z=LOAD *A
3073	
3074	in that order, but, without intervention, the sequence may have almost any
3075	combination of elements combined or discarded, provided the program's view
3076	of the world remains consistent.  Note that READ_ONCE() and WRITE_ONCE()
3077	are -not- optional in the above example, as there are architectures
3078	where a given CPU might reorder successive loads to the same location.
3079	On such architectures, READ_ONCE() and WRITE_ONCE() do whatever is
3080	necessary to prevent this, for example, on Itanium the volatile casts
3081	used by READ_ONCE() and WRITE_ONCE() cause GCC to emit the special ld.acq
3082	and st.rel instructions (respectively) that prevent such reordering.
3083	
3084	The compiler may also combine, discard or defer elements of the sequence before
3085	the CPU even sees them.
3086	
3087	For instance:
3088	
3089		*A = V;
3090		*A = W;
3091	
3092	may be reduced to:
3093	
3094		*A = W;
3095	
3096	since, without either a write barrier or an WRITE_ONCE(), it can be
3097	assumed that the effect of the storage of V to *A is lost.  Similarly:
3098	
3099		*A = Y;
3100		Z = *A;
3101	
3102	may, without a memory barrier or an READ_ONCE() and WRITE_ONCE(), be
3103	reduced to:
3104	
3105		*A = Y;
3106		Z = Y;
3107	
3108	and the LOAD operation never appear outside of the CPU.
3109	
3110	
3111	AND THEN THERE'S THE ALPHA
3112	--------------------------
3113	
3114	The DEC Alpha CPU is one of the most relaxed CPUs there is.  Not only that,
3115	some versions of the Alpha CPU have a split data cache, permitting them to have
3116	two semantically-related cache lines updated at separate times.  This is where
3117	the data dependency barrier really becomes necessary as this synchronises both
3118	caches with the memory coherence system, thus making it seem like pointer
3119	changes vs new data occur in the right order.
3120	
3121	The Alpha defines the Linux kernel's memory barrier model.
3122	
3123	See the subsection on "Cache Coherency" above.
3124	
3125	
3126	VIRTUAL MACHINE GUESTS
3127	----------------------
3128	
3129	Guests running within virtual machines might be affected by SMP effects even if
3130	the guest itself is compiled without SMP support.  This is an artifact of
3131	interfacing with an SMP host while running an UP kernel.  Using mandatory
3132	barriers for this use-case would be possible but is often suboptimal.
3133	
3134	To handle this case optimally, low-level virt_mb() etc macros are available.
3135	These have the same effect as smp_mb() etc when SMP is enabled, but generate
3136	identical code for SMP and non-SMP systems.  For example, virtual machine guests
3137	should use virt_mb() rather than smp_mb() when synchronizing against a
3138	(possibly SMP) host.
3139	
3140	These are equivalent to smp_mb() etc counterparts in all other respects,
3141	in particular, they do not control MMIO effects: to control
3142	MMIO effects, use mandatory barriers.
3143	
3144	
3145	============
3146	EXAMPLE USES
3147	============
3148	
3149	CIRCULAR BUFFERS
3150	----------------
3151	
3152	Memory barriers can be used to implement circular buffering without the need
3153	of a lock to serialise the producer with the consumer.  See:
3154	
3155		Documentation/circular-buffers.txt
3156	
3157	for details.
3158	
3159	
3160	==========
3161	REFERENCES
3162	==========
3163	
3164	Alpha AXP Architecture Reference Manual, Second Edition (Sites & Witek,
3165	Digital Press)
3166		Chapter 5.2: Physical Address Space Characteristics
3167		Chapter 5.4: Caches and Write Buffers
3168		Chapter 5.5: Data Sharing
3169		Chapter 5.6: Read/Write Ordering
3170	
3171	AMD64 Architecture Programmer's Manual Volume 2: System Programming
3172		Chapter 7.1: Memory-Access Ordering
3173		Chapter 7.4: Buffering and Combining Memory Writes
3174	
3175	IA-32 Intel Architecture Software Developer's Manual, Volume 3:
3176	System Programming Guide
3177		Chapter 7.1: Locked Atomic Operations
3178		Chapter 7.2: Memory Ordering
3179		Chapter 7.4: Serializing Instructions
3180	
3181	The SPARC Architecture Manual, Version 9
3182		Chapter 8: Memory Models
3183		Appendix D: Formal Specification of the Memory Models
3184		Appendix J: Programming with the Memory Models
3185	
3186	UltraSPARC Programmer Reference Manual
3187		Chapter 5: Memory Accesses and Cacheability
3188		Chapter 15: Sparc-V9 Memory Models
3189	
3190	UltraSPARC III Cu User's Manual
3191		Chapter 9: Memory Models
3192	
3193	UltraSPARC IIIi Processor User's Manual
3194		Chapter 8: Memory Models
3195	
3196	UltraSPARC Architecture 2005
3197		Chapter 9: Memory
3198		Appendix D: Formal Specifications of the Memory Models
3199	
3200	UltraSPARC T1 Supplement to the UltraSPARC Architecture 2005
3201		Chapter 8: Memory Models
3202		Appendix F: Caches and Cache Coherency
3203	
3204	Solaris Internals, Core Kernel Architecture, p63-68:
3205		Chapter 3.3: Hardware Considerations for Locks and
3206				Synchronization
3207	
3208	Unix Systems for Modern Architectures, Symmetric Multiprocessing and Caching
3209	for Kernel Programmers:
3210		Chapter 13: Other Memory Models
3211	
3212	Intel Itanium Architecture Software Developer's Manual: Volume 1:
3213		Section 2.6: Speculation
3214		Section 4.4: Memory Access
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