Adding upstream version 6.6.15.upstream/6.6.15

Signed-off-by: Daniel Baumann <daniel.baumann@progress-linux.org>

71 files changed, 10235 insertions, 0 deletions

diff --git a/tools/memory-model/.gitignore b/tools/memory-model/.gitignore
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+# SPDX-License-Identifier: GPL-2.0-only
+litmus
diff --git a/tools/memory-model/Documentation/README b/tools/memory-model/Documentation/README
new file mode 100644
index 0000000000..db90a26dbd
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+++ b/tools/memory-model/Documentation/README
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+It has been said that successful communication requires first identifying
+what your audience knows and then building a bridge from their current
+knowledge to what they need to know.  Unfortunately, the expected
+Linux-kernel memory model (LKMM) audience might be anywhere from novice
+to expert both in kernel hacking and in understanding LKMM.
+
+This document therefore points out a number of places to start reading,
+depending on what you know and what you would like to learn.  Please note
+that the documents later in this list assume that the reader understands
+the material provided by documents earlier in this list.
+
+o	You are new to Linux-kernel concurrency: simple.txt
+
+o	You have some background in Linux-kernel concurrency, and would
+	like an overview of the types of low-level concurrency primitives
+	that the Linux kernel provides:  ordering.txt
+
+	Here, "low level" means atomic operations to single variables.
+
+o	You are familiar with the Linux-kernel concurrency primitives
+	that you need, and just want to get started with LKMM litmus
+	tests:  litmus-tests.txt
+
+o	You are familiar with Linux-kernel concurrency, and would
+	like a detailed intuitive understanding of LKMM, including
+	situations involving more than two threads:  recipes.txt
+
+o	You would like a detailed understanding of what your compiler can
+	and cannot do to control dependencies:  control-dependencies.txt
+
+o	You are familiar with Linux-kernel concurrency and the use of
+	LKMM, and would like a quick reference:  cheatsheet.txt
+
+o	You are familiar with Linux-kernel concurrency and the use
+	of LKMM, and would like to learn about LKMM's requirements,
+	rationale, and implementation:	explanation.txt
+
+o	You are interested in the publications related to LKMM, including
+	hardware manuals, academic literature, standards-committee
+	working papers, and LWN articles:  references.txt
+
+
+====================
+DESCRIPTION OF FILES
+====================
+
+README
+	This file.
+
+cheatsheet.txt
+	Quick-reference guide to the Linux-kernel memory model.
+
+control-dependencies.txt
+	Guide to preventing compiler optimizations from destroying
+	your control dependencies.
+
+explanation.txt
+	Detailed description of the memory model.
+
+litmus-tests.txt
+	The format, features, capabilities, and limitations of the litmus
+	tests that LKMM can evaluate.
+
+ordering.txt
+	Overview of the Linux kernel's low-level memory-ordering
+	primitives by category.
+
+recipes.txt
+	Common memory-ordering patterns.
+
+references.txt
+	Background information.
+
+simple.txt
+	Starting point for someone new to Linux-kernel concurrency.
+	And also a reminder of the simpler approaches to concurrency!
diff --git a/tools/memory-model/Documentation/access-marking.txt b/tools/memory-model/Documentation/access-marking.txt
new file mode 100644
index 0000000000..6577822218
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+++ b/tools/memory-model/Documentation/access-marking.txt
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+MARKING SHARED-MEMORY ACCESSES
+==============================
+
+This document provides guidelines for marking intentionally concurrent
+normal accesses to shared memory, that is "normal" as in accesses that do
+not use read-modify-write atomic operations.  It also describes how to
+document these accesses, both with comments and with special assertions
+processed by the Kernel Concurrency Sanitizer (KCSAN).  This discussion
+builds on an earlier LWN article [1].
+
+
+ACCESS-MARKING OPTIONS
+======================
+
+The Linux kernel provides the following access-marking options:
+
+1.	Plain C-language accesses (unmarked), for example, "a = b;"
+
+2.	Data-race marking, for example, "data_race(a = b);"
+
+3.	READ_ONCE(), for example, "a = READ_ONCE(b);"
+	The various forms of atomic_read() also fit in here.
+
+4.	WRITE_ONCE(), for example, "WRITE_ONCE(a, b);"
+	The various forms of atomic_set() also fit in here.
+
+
+These may be used in combination, as shown in this admittedly improbable
+example:
+
+	WRITE_ONCE(a, b + data_race(c + d) + READ_ONCE(e));
+
+Neither plain C-language accesses nor data_race() (#1 and #2 above) place
+any sort of constraint on the compiler's choice of optimizations [2].
+In contrast, READ_ONCE() and WRITE_ONCE() (#3 and #4 above) restrict the
+compiler's use of code-motion and common-subexpression optimizations.
+Therefore, if a given access is involved in an intentional data race,
+using READ_ONCE() for loads and WRITE_ONCE() for stores is usually
+preferable to data_race(), which in turn is usually preferable to plain
+C-language accesses.  It is permissible to combine #2 and #3, for example,
+data_race(READ_ONCE(a)), which will both restrict compiler optimizations
+and disable KCSAN diagnostics.
+
+KCSAN will complain about many types of data races involving plain
+C-language accesses, but marking all accesses involved in a given data
+race with one of data_race(), READ_ONCE(), or WRITE_ONCE(), will prevent
+KCSAN from complaining.  Of course, lack of KCSAN complaints does not
+imply correct code.  Therefore, please take a thoughtful approach
+when responding to KCSAN complaints.  Churning the code base with
+ill-considered additions of data_race(), READ_ONCE(), and WRITE_ONCE()
+is unhelpful.
+
+In fact, the following sections describe situations where use of
+data_race() and even plain C-language accesses is preferable to
+READ_ONCE() and WRITE_ONCE().
+
+
+Use of the data_race() Macro
+----------------------------
+
+Here are some situations where data_race() should be used instead of
+READ_ONCE() and WRITE_ONCE():
+
+1.	Data-racy loads from shared variables whose values are used only
+	for diagnostic purposes.
+
+2.	Data-racy reads whose values are checked against marked reload.
+
+3.	Reads whose values feed into error-tolerant heuristics.
+
+4.	Writes setting values that feed into error-tolerant heuristics.
+
+
+Data-Racy Reads for Approximate Diagnostics
+
+Approximate diagnostics include lockdep reports, monitoring/statistics
+(including /proc and /sys output), WARN*()/BUG*() checks whose return
+values are ignored, and other situations where reads from shared variables
+are not an integral part of the core concurrency design.
+
+In fact, use of data_race() instead READ_ONCE() for these diagnostic
+reads can enable better checking of the remaining accesses implementing
+the core concurrency design.  For example, suppose that the core design
+prevents any non-diagnostic reads from shared variable x from running
+concurrently with updates to x.  Then using plain C-language writes
+to x allows KCSAN to detect reads from x from within regions of code
+that fail to exclude the updates.  In this case, it is important to use
+data_race() for the diagnostic reads because otherwise KCSAN would give
+false-positive warnings about these diagnostic reads.
+
+If it is necessary to both restrict compiler optimizations and disable
+KCSAN diagnostics, use both data_race() and READ_ONCE(), for example,
+data_race(READ_ONCE(a)).
+
+In theory, plain C-language loads can also be used for this use case.
+However, in practice this will have the disadvantage of causing KCSAN
+to generate false positives because KCSAN will have no way of knowing
+that the resulting data race was intentional.
+
+
+Data-Racy Reads That Are Checked Against Marked Reload
+
+The values from some reads are not implicitly trusted.  They are instead
+fed into some operation that checks the full value against a later marked
+load from memory, which means that the occasional arbitrarily bogus value
+is not a problem.  For example, if a bogus value is fed into cmpxchg(),
+all that happens is that this cmpxchg() fails, which normally results
+in a retry.  Unless the race condition that resulted in the bogus value
+recurs, this retry will with high probability succeed, so no harm done.
+
+However, please keep in mind that a data_race() load feeding into
+a cmpxchg_relaxed() might still be subject to load fusing on some
+architectures.  Therefore, it is best to capture the return value from
+the failing cmpxchg() for the next iteration of the loop, an approach
+that provides the compiler much less scope for mischievous optimizations.
+Capturing the return value from cmpxchg() also saves a memory reference
+in many cases.
+
+In theory, plain C-language loads can also be used for this use case.
+However, in practice this will have the disadvantage of causing KCSAN
+to generate false positives because KCSAN will have no way of knowing
+that the resulting data race was intentional.
+
+
+Reads Feeding Into Error-Tolerant Heuristics
+
+Values from some reads feed into heuristics that can tolerate occasional
+errors.  Such reads can use data_race(), thus allowing KCSAN to focus on
+the other accesses to the relevant shared variables.  But please note
+that data_race() loads are subject to load fusing, which can result in
+consistent errors, which in turn are quite capable of breaking heuristics.
+Therefore use of data_race() should be limited to cases where some other
+code (such as a barrier() call) will force the occasional reload.
+
+Note that this use case requires that the heuristic be able to handle
+any possible error.  In contrast, if the heuristics might be fatally
+confused by one or more of the possible erroneous values, use READ_ONCE()
+instead of data_race().
+
+In theory, plain C-language loads can also be used for this use case.
+However, in practice this will have the disadvantage of causing KCSAN
+to generate false positives because KCSAN will have no way of knowing
+that the resulting data race was intentional.
+
+
+Writes Setting Values Feeding Into Error-Tolerant Heuristics
+
+The values read into error-tolerant heuristics come from somewhere,
+for example, from sysfs.  This means that some code in sysfs writes
+to this same variable, and these writes can also use data_race().
+After all, if the heuristic can tolerate the occasional bogus value
+due to compiler-mangled reads, it can also tolerate the occasional
+compiler-mangled write, at least assuming that the proper value is in
+place once the write completes.
+
+Plain C-language stores can also be used for this use case.  However,
+in kernels built with CONFIG_KCSAN_ASSUME_PLAIN_WRITES_ATOMIC=n, this
+will have the disadvantage of causing KCSAN to generate false positives
+because KCSAN will have no way of knowing that the resulting data race
+was intentional.
+
+
+Use of Plain C-Language Accesses
+--------------------------------
+
+Here are some example situations where plain C-language accesses should
+used instead of READ_ONCE(), WRITE_ONCE(), and data_race():
+
+1.	Accesses protected by mutual exclusion, including strict locking
+	and sequence locking.
+
+2.	Initialization-time and cleanup-time accesses.	This covers a
+	wide variety of situations, including the uniprocessor phase of
+	system boot, variables to be used by not-yet-spawned kthreads,
+	structures not yet published to reference-counted or RCU-protected
+	data structures, and the cleanup side of any of these situations.
+
+3.	Per-CPU variables that are not accessed from other CPUs.
+
+4.	Private per-task variables, including on-stack variables, some
+	fields in the task_struct structure, and task-private heap data.
+
+5.	Any other loads for which there is not supposed to be a concurrent
+	store to that same variable.
+
+6.	Any other stores for which there should be neither concurrent
+	loads nor concurrent stores to that same variable.
+
+	But note that KCSAN makes two explicit exceptions to this rule
+	by default, refraining from flagging plain C-language stores:
+
+	a.	No matter what.  You can override this default by building
+		with CONFIG_KCSAN_ASSUME_PLAIN_WRITES_ATOMIC=n.
+
+	b.	When the store writes the value already contained in
+		that variable.	You can override this default by building
+		with CONFIG_KCSAN_REPORT_VALUE_CHANGE_ONLY=n.
+
+	c.	When one of the stores is in an interrupt handler and
+		the other in the interrupted code.  You can override this
+		default by building with CONFIG_KCSAN_INTERRUPT_WATCHER=y.
+
+Note that it is important to use plain C-language accesses in these cases,
+because doing otherwise prevents KCSAN from detecting violations of your
+code's synchronization rules.
+
+
+ACCESS-DOCUMENTATION OPTIONS
+============================
+
+It is important to comment marked accesses so that people reading your
+code, yourself included, are reminded of the synchronization design.
+However, it is even more important to comment plain C-language accesses
+that are intentionally involved in data races.  Such comments are
+needed to remind people reading your code, again, yourself included,
+of how the compiler has been prevented from optimizing those accesses
+into concurrency bugs.
+
+It is also possible to tell KCSAN about your synchronization design.
+For example, ASSERT_EXCLUSIVE_ACCESS(foo) tells KCSAN that any
+concurrent access to variable foo by any other CPU is an error, even
+if that concurrent access is marked with READ_ONCE().  In addition,
+ASSERT_EXCLUSIVE_WRITER(foo) tells KCSAN that although it is OK for there
+to be concurrent reads from foo from other CPUs, it is an error for some
+other CPU to be concurrently writing to foo, even if that concurrent
+write is marked with data_race() or WRITE_ONCE().
+
+Note that although KCSAN will call out data races involving either
+ASSERT_EXCLUSIVE_ACCESS() or ASSERT_EXCLUSIVE_WRITER() on the one hand
+and data_race() writes on the other, KCSAN will not report the location
+of these data_race() writes.
+
+
+EXAMPLES
+========
+
+As noted earlier, the goal is to prevent the compiler from destroying
+your concurrent algorithm, to help the human reader, and to inform
+KCSAN of aspects of your concurrency design.  This section looks at a
+few examples showing how this can be done.
+
+
+Lock Protection With Lockless Diagnostic Access
+-----------------------------------------------
+
+For example, suppose a shared variable "foo" is read only while a
+reader-writer spinlock is read-held, written only while that same
+spinlock is write-held, except that it is also read locklessly for
+diagnostic purposes.  The code might look as follows:
+
+	int foo;
+	DEFINE_RWLOCK(foo_rwlock);
+
+	void update_foo(int newval)
+	{
+		write_lock(&foo_rwlock);
+		foo = newval;
+		do_something(newval);
+		write_unlock(&foo_rwlock);
+	}
+
+	int read_foo(void)
+	{
+		int ret;
+
+		read_lock(&foo_rwlock);
+		do_something_else();
+		ret = foo;
+		read_unlock(&foo_rwlock);
+		return ret;
+	}
+
+	void read_foo_diagnostic(void)
+	{
+		pr_info("Current value of foo: %d\n", data_race(foo));
+	}
+
+The reader-writer lock prevents the compiler from introducing concurrency
+bugs into any part of the main algorithm using foo, which means that
+the accesses to foo within both update_foo() and read_foo() can (and
+should) be plain C-language accesses.  One benefit of making them be
+plain C-language accesses is that KCSAN can detect any erroneous lockless
+reads from or updates to foo.  The data_race() in read_foo_diagnostic()
+tells KCSAN that data races are expected, and should be silently
+ignored.  This data_race() also tells the human reading the code that
+read_foo_diagnostic() might sometimes return a bogus value.
+
+If it is necessary to suppress compiler optimization and also detect
+buggy lockless writes, read_foo_diagnostic() can be updated as follows:
+
+	void read_foo_diagnostic(void)
+	{
+		pr_info("Current value of foo: %d\n", data_race(READ_ONCE(foo)));
+	}
+
+Alternatively, given that KCSAN is to ignore all accesses in this function,
+this function can be marked __no_kcsan and the data_race() can be dropped:
+
+	void __no_kcsan read_foo_diagnostic(void)
+	{
+		pr_info("Current value of foo: %d\n", READ_ONCE(foo));
+	}
+
+However, in order for KCSAN to detect buggy lockless writes, your kernel
+must be built with CONFIG_KCSAN_ASSUME_PLAIN_WRITES_ATOMIC=n.  If you
+need KCSAN to detect such a write even if that write did not change
+the value of foo, you also need CONFIG_KCSAN_REPORT_VALUE_CHANGE_ONLY=n.
+If you need KCSAN to detect such a write happening in an interrupt handler
+running on the same CPU doing the legitimate lock-protected write, you
+also need CONFIG_KCSAN_INTERRUPT_WATCHER=y.  With some or all of these
+Kconfig options set properly, KCSAN can be quite helpful, although
+it is not necessarily a full replacement for hardware watchpoints.
+On the other hand, neither are hardware watchpoints a full replacement
+for KCSAN because it is not always easy to tell hardware watchpoint to
+conditionally trap on accesses.
+
+
+Lock-Protected Writes With Lockless Reads
+-----------------------------------------
+
+For another example, suppose a shared variable "foo" is updated only
+while holding a spinlock, but is read locklessly.  The code might look
+as follows:
+
+	int foo;
+	DEFINE_SPINLOCK(foo_lock);
+
+	void update_foo(int newval)
+	{
+		spin_lock(&foo_lock);
+		WRITE_ONCE(foo, newval);
+		ASSERT_EXCLUSIVE_WRITER(foo);
+		do_something(newval);
+		spin_unlock(&foo_wlock);
+	}
+
+	int read_foo(void)
+	{
+		do_something_else();
+		return READ_ONCE(foo);
+	}
+
+Because foo is read locklessly, all accesses are marked.  The purpose
+of the ASSERT_EXCLUSIVE_WRITER() is to allow KCSAN to check for a buggy
+concurrent lockless write.
+
+
+Lock-Protected Writes With Heuristic Lockless Reads
+---------------------------------------------------
+
+For another example, suppose that the code can normally make use of
+a per-data-structure lock, but there are times when a global lock
+is required.  These times are indicated via a global flag.  The code
+might look as follows, and is based loosely on nf_conntrack_lock(),
+nf_conntrack_all_lock(), and nf_conntrack_all_unlock():
+
+	bool global_flag;
+	DEFINE_SPINLOCK(global_lock);
+	struct foo {
+		spinlock_t f_lock;
+		int f_data;
+	};
+
+	/* All foo structures are in the following array. */
+	int nfoo;
+	struct foo *foo_array;
+
+	void do_something_locked(struct foo *fp)
+	{
+		/* This works even if data_race() returns nonsense. */
+		if (!data_race(global_flag)) {
+			spin_lock(&fp->f_lock);
+			if (!smp_load_acquire(&global_flag)) {
+				do_something(fp);
+				spin_unlock(&fp->f_lock);
+				return;
+			}
+			spin_unlock(&fp->f_lock);
+		}
+		spin_lock(&global_lock);
+		/* global_lock held, thus global flag cannot be set. */
+		spin_lock(&fp->f_lock);
+		spin_unlock(&global_lock);
+		/*
+		 * global_flag might be set here, but begin_global()
+		 * will wait for ->f_lock to be released.
+		 */
+		do_something(fp);
+		spin_unlock(&fp->f_lock);
+	}
+
+	void begin_global(void)
+	{
+		int i;
+
+		spin_lock(&global_lock);
+		WRITE_ONCE(global_flag, true);
+		for (i = 0; i < nfoo; i++) {
+			/*
+			 * Wait for pre-existing local locks.  One at
+			 * a time to avoid lockdep limitations.
+			 */
+			spin_lock(&fp->f_lock);
+			spin_unlock(&fp->f_lock);
+		}
+	}
+
+	void end_global(void)
+	{
+		smp_store_release(&global_flag, false);
+		spin_unlock(&global_lock);
+	}
+
+All code paths leading from the do_something_locked() function's first
+read from global_flag acquire a lock, so endless load fusing cannot
+happen.
+
+If the value read from global_flag is true, then global_flag is
+rechecked while holding ->f_lock, which, if global_flag is now false,
+prevents begin_global() from completing.  It is therefore safe to invoke
+do_something().
+
+Otherwise, if either value read from global_flag is true, then after
+global_lock is acquired global_flag must be false.  The acquisition of
+->f_lock will prevent any call to begin_global() from returning, which
+means that it is safe to release global_lock and invoke do_something().
+
+For this to work, only those foo structures in foo_array[] may be passed
+to do_something_locked().  The reason for this is that the synchronization
+with begin_global() relies on momentarily holding the lock of each and
+every foo structure.
+
+The smp_load_acquire() and smp_store_release() are required because
+changes to a foo structure between calls to begin_global() and
+end_global() are carried out without holding that structure's ->f_lock.
+The smp_load_acquire() and smp_store_release() ensure that the next
+invocation of do_something() from do_something_locked() will see those
+changes.
+
+
+Lockless Reads and Writes
+-------------------------
+
+For another example, suppose a shared variable "foo" is both read and
+updated locklessly.  The code might look as follows:
+
+	int foo;
+
+	int update_foo(int newval)
+	{
+		int ret;
+
+		ret = xchg(&foo, newval);
+		do_something(newval);
+		return ret;
+	}
+
+	int read_foo(void)
+	{
+		do_something_else();
+		return READ_ONCE(foo);
+	}
+
+Because foo is accessed locklessly, all accesses are marked.  It does
+not make sense to use ASSERT_EXCLUSIVE_WRITER() in this case because
+there really can be concurrent lockless writers.  KCSAN would
+flag any concurrent plain C-language reads from foo, and given
+CONFIG_KCSAN_ASSUME_PLAIN_WRITES_ATOMIC=n, also any concurrent plain
+C-language writes to foo.
+
+
+Lockless Reads and Writes, But With Single-Threaded Initialization
+------------------------------------------------------------------
+
+For yet another example, suppose that foo is initialized in a
+single-threaded manner, but that a number of kthreads are then created
+that locklessly and concurrently access foo.  Some snippets of this code
+might look as follows:
+
+	int foo;
+
+	void initialize_foo(int initval, int nkthreads)
+	{
+		int i;
+
+		foo = initval;
+		ASSERT_EXCLUSIVE_ACCESS(foo);
+		for (i = 0; i < nkthreads; i++)
+			kthread_run(access_foo_concurrently, ...);
+	}
+
+	/* Called from access_foo_concurrently(). */
+	int update_foo(int newval)
+	{
+		int ret;
+
+		ret = xchg(&foo, newval);
+		do_something(newval);
+		return ret;
+	}
+
+	/* Also called from access_foo_concurrently(). */
+	int read_foo(void)
+	{
+		do_something_else();
+		return READ_ONCE(foo);
+	}
+
+The initialize_foo() uses a plain C-language write to foo because there
+are not supposed to be concurrent accesses during initialization.  The
+ASSERT_EXCLUSIVE_ACCESS() allows KCSAN to flag buggy concurrent unmarked
+reads, and the ASSERT_EXCLUSIVE_ACCESS() call further allows KCSAN to
+flag buggy concurrent writes, even if:  (1) Those writes are marked or
+(2) The kernel was built with CONFIG_KCSAN_ASSUME_PLAIN_WRITES_ATOMIC=y.
+
+
+Checking Stress-Test Race Coverage
+----------------------------------
+
+When designing stress tests it is important to ensure that race conditions
+of interest really do occur.  For example, consider the following code
+fragment:
+
+	int foo;
+
+	int update_foo(int newval)
+	{
+		return xchg(&foo, newval);
+	}
+
+	int xor_shift_foo(int shift, int mask)
+	{
+		int old, new, newold;
+
+		newold = data_race(foo); /* Checked by cmpxchg(). */
+		do {
+			old = newold;
+			new = (old << shift) ^ mask;
+			newold = cmpxchg(&foo, old, new);
+		} while (newold != old);
+		return old;
+	}
+
+	int read_foo(void)
+	{
+		return READ_ONCE(foo);
+	}
+
+If it is possible for update_foo(), xor_shift_foo(), and read_foo() to be
+invoked concurrently, the stress test should force this concurrency to
+actually happen.  KCSAN can evaluate the stress test when the above code
+is modified to read as follows:
+
+	int foo;
+
+	int update_foo(int newval)
+	{
+		ASSERT_EXCLUSIVE_ACCESS(foo);
+		return xchg(&foo, newval);
+	}
+
+	int xor_shift_foo(int shift, int mask)
+	{
+		int old, new, newold;
+
+		newold = data_race(foo); /* Checked by cmpxchg(). */
+		do {
+			old = newold;
+			new = (old << shift) ^ mask;
+			ASSERT_EXCLUSIVE_ACCESS(foo);
+			newold = cmpxchg(&foo, old, new);
+		} while (newold != old);
+		return old;
+	}
+
+
+	int read_foo(void)
+	{
+		ASSERT_EXCLUSIVE_ACCESS(foo);
+		return READ_ONCE(foo);
+	}
+
+If a given stress-test run does not result in KCSAN complaints from
+each possible pair of ASSERT_EXCLUSIVE_ACCESS() invocations, the
+stress test needs improvement.  If the stress test was to be evaluated
+on a regular basis, it would be wise to place the above instances of
+ASSERT_EXCLUSIVE_ACCESS() under #ifdef so that they did not result in
+false positives when not evaluating the stress test.
+
+
+REFERENCES
+==========
+
+[1] "Concurrency bugs should fear the big bad data-race detector (part 2)"
+    https://lwn.net/Articles/816854/
+
+[2] "Who's afraid of a big bad optimizing compiler?"
+    https://lwn.net/Articles/793253/
diff --git a/tools/memory-model/Documentation/cheatsheet.txt b/tools/memory-model/Documentation/cheatsheet.txt
new file mode 100644
index 0000000000..99d00870b1
--- /dev/null
+++ b/tools/memory-model/Documentation/cheatsheet.txt
@@ -0,0 +1,35 @@
+                                  Prior Operation     Subsequent Operation
+                                  ---------------  ---------------------------
+                               C  Self  R  W  RMW  Self  R  W  DR  DW  RMW  SV
+                              --  ----  -  -  ---  ----  -  -  --  --  ---  --
+
+Relaxed store                        Y                                       Y
+Relaxed load                         Y                          Y   Y        Y
+Relaxed RMW operation                Y                          Y   Y        Y
+rcu_dereference()                    Y                          Y   Y        Y
+Successful *_acquire()               R                   Y  Y   Y   Y    Y   Y
+Successful *_release()         C        Y  Y    Y     W                      Y
+smp_rmb()                               Y       R        Y      Y        R
+smp_wmb()                                  Y    W           Y       Y    W
+smp_mb() & synchronize_rcu()  CP        Y  Y    Y        Y  Y   Y   Y    Y
+Successful full non-void RMW  CP     Y  Y  Y    Y     Y  Y  Y   Y   Y    Y   Y
+smp_mb__before_atomic()       CP        Y  Y    Y        a  a   a   a    Y
+smp_mb__after_atomic()        CP        a  a    Y        Y  Y   Y   Y    Y
+
+
+Key:	Relaxed:  A relaxed operation is either READ_ONCE(), WRITE_ONCE(),
+		  a *_relaxed() RMW operation, an unsuccessful RMW
+		  operation, a non-value-returning RMW operation such
+		  as atomic_inc(), or one of the atomic*_read() and
+		  atomic*_set() family of operations.
+	C:	  Ordering is cumulative
+	P:	  Ordering propagates
+	R:	  Read, for example, READ_ONCE(), or read portion of RMW
+	W:	  Write, for example, WRITE_ONCE(), or write portion of RMW
+	Y:	  Provides ordering
+	a:	  Provides ordering given intervening RMW atomic operation
+	DR:	  Dependent read (address dependency)
+	DW:	  Dependent write (address, data, or control dependency)
+	RMW:	  Atomic read-modify-write operation
+	SELF:	  Orders self, as opposed to accesses before and/or after
+	SV:	  Orders later accesses to the same variable
diff --git a/tools/memory-model/Documentation/control-dependencies.txt b/tools/memory-model/Documentation/control-dependencies.txt
new file mode 100644
index 0000000000..8b743d20fe
--- /dev/null
+++ b/tools/memory-model/Documentation/control-dependencies.txt
@@ -0,0 +1,258 @@
+CONTROL DEPENDENCIES
+====================
+
+A major difficulty with control dependencies is that current compilers
+do not support them.  One purpose of this document is therefore to
+help you prevent your compiler from breaking your code.  However,
+control dependencies also pose other challenges, which leads to the
+second purpose of this document, namely to help you to avoid breaking
+your own code, even in the absence of help from your compiler.
+
+One such challenge is that control dependencies order only later stores.
+Therefore, a load-load control dependency will not preserve ordering
+unless a read memory barrier is provided.  Consider the following code:
+
+	q = READ_ONCE(a);
+	if (q)
+		p = READ_ONCE(b);
+
+This is not guaranteed to provide any ordering because some types of CPUs
+are permitted to predict the result of the load from "b".  This prediction
+can cause other CPUs to see this load as having happened before the load
+from "a".  This means that an explicit read barrier is required, for example
+as follows:
+
+	q = READ_ONCE(a);
+	if (q) {
+		smp_rmb();
+		p = READ_ONCE(b);
+	}
+
+However, stores are not speculated.  This means that ordering is
+(usually) guaranteed for load-store control dependencies, as in the
+following example:
+
+	q = READ_ONCE(a);
+	if (q)
+		WRITE_ONCE(b, 1);
+
+Control dependencies can pair with each other and with other types
+of ordering.  But please note that neither the READ_ONCE() nor the
+WRITE_ONCE() are optional.  Without the READ_ONCE(), the compiler might
+fuse the load from "a" with other loads.  Without the WRITE_ONCE(),
+the compiler might fuse the store to "b" with other stores.  Worse yet,
+the compiler might convert the store into a load and a check followed
+by a store, and this compiler-generated load would not be ordered by
+the control dependency.
+
+Furthermore, if the compiler is able to prove that the value of variable
+"a" is always non-zero, it would be well within its rights to optimize
+the original example by eliminating the "if" statement as follows:
+
+	q = a;
+	b = 1;  /* BUG: Compiler and CPU can both reorder!!! */
+
+So don't leave out either the READ_ONCE() or the WRITE_ONCE().
+In particular, although READ_ONCE() does force the compiler to emit a
+load, it does *not* force the compiler to actually use the loaded value.
+
+It is tempting to try use control dependencies to enforce ordering on
+identical stores on both branches of the "if" statement as follows:
+
+	q = READ_ONCE(a);
+	if (q) {
+		barrier();
+		WRITE_ONCE(b, 1);
+		do_something();
+	} else {
+		barrier();
+		WRITE_ONCE(b, 1);
+		do_something_else();
+	}
+
+Unfortunately, current compilers will transform this as follows at high
+optimization levels:
+
+	q = READ_ONCE(a);
+	barrier();
+	WRITE_ONCE(b, 1);  /* BUG: No ordering vs. load from a!!! */
+	if (q) {
+		/* WRITE_ONCE(b, 1); -- moved up, BUG!!! */
+		do_something();
+	} else {
+		/* WRITE_ONCE(b, 1); -- moved up, BUG!!! */
+		do_something_else();
+	}
+
+Now there is no conditional between the load from "a" and the store to
+"b", which means that the CPU is within its rights to reorder them:  The
+conditional is absolutely required, and must be present in the final
+assembly code, after all of the compiler and link-time optimizations
+have been applied.  Therefore, if you need ordering in this example,
+you must use explicit memory ordering, for example, smp_store_release():
+
+	q = READ_ONCE(a);
+	if (q) {
+		smp_store_release(&b, 1);
+		do_something();
+	} else {
+		smp_store_release(&b, 1);
+		do_something_else();
+	}
+
+Without explicit memory ordering, control-dependency-based ordering is
+guaranteed only when the stores differ, for example:
+
+	q = READ_ONCE(a);
+	if (q) {
+		WRITE_ONCE(b, 1);
+		do_something();
+	} else {
+		WRITE_ONCE(b, 2);
+		do_something_else();
+	}
+
+The initial READ_ONCE() is still required to prevent the compiler from
+knowing too much about the value of "a".
+
+But please note that you need to be careful what you do with the local
+variable "q", otherwise the compiler might be able to guess the value
+and again remove the conditional branch that is absolutely required to
+preserve ordering.  For example:
+
+	q = READ_ONCE(a);
+	if (q % MAX) {
+		WRITE_ONCE(b, 1);
+		do_something();
+	} else {
+		WRITE_ONCE(b, 2);
+		do_something_else();
+	}
+
+If MAX is compile-time defined to be 1, then the compiler knows that
+(q % MAX) must be equal to zero, regardless of the value of "q".
+The compiler is therefore within its rights to transform the above code
+into the following:
+
+	q = READ_ONCE(a);
+	WRITE_ONCE(b, 2);
+	do_something_else();
+
+Given this transformation, the CPU is not required to respect the ordering
+between the load from variable "a" and the store to variable "b".  It is
+tempting to add a barrier(), but this does not help.  The conditional
+is gone, and the barrier won't bring it back.  Therefore, if you need
+to relying on control dependencies to produce this ordering, you should
+make sure that MAX is greater than one, perhaps as follows:
+
+	q = READ_ONCE(a);
+	BUILD_BUG_ON(MAX <= 1); /* Order load from a with store to b. */
+	if (q % MAX) {
+		WRITE_ONCE(b, 1);
+		do_something();
+	} else {
+		WRITE_ONCE(b, 2);
+		do_something_else();
+	}
+
+Please note once again that each leg of the "if" statement absolutely
+must store different values to "b".  As in previous examples, if the two
+values were identical, the compiler could pull this store outside of the
+"if" statement, destroying the control dependency's ordering properties.
+
+You must also be careful avoid relying too much on boolean short-circuit
+evaluation.  Consider this example:
+
+	q = READ_ONCE(a);
+	if (q || 1 > 0)
+		WRITE_ONCE(b, 1);
+
+Because the first condition cannot fault and the second condition is
+always true, the compiler can transform this example as follows, again
+destroying the control dependency's ordering:
+
+	q = READ_ONCE(a);
+	WRITE_ONCE(b, 1);
+
+This is yet another example showing the importance of preventing the
+compiler from out-guessing your code.  Again, although READ_ONCE() really
+does force the compiler to emit code for a given load, the compiler is
+within its rights to discard the loaded value.
+
+In addition, control dependencies apply only to the then-clause and
+else-clause of the "if" statement in question.  In particular, they do
+not necessarily order the code following the entire "if" statement:
+
+	q = READ_ONCE(a);
+	if (q) {
+		WRITE_ONCE(b, 1);
+	} else {
+		WRITE_ONCE(b, 2);
+	}
+	WRITE_ONCE(c, 1);  /* BUG: No ordering against the read from "a". */
+
+It is tempting to argue that there in fact is ordering because the
+compiler cannot reorder volatile accesses and also cannot reorder
+the writes to "b" with the condition.  Unfortunately for this line
+of reasoning, the compiler might compile the two writes to "b" as
+conditional-move instructions, as in this fanciful pseudo-assembly
+language:
+
+	ld r1,a
+	cmp r1,$0
+	cmov,ne r4,$1
+	cmov,eq r4,$2
+	st r4,b
+	st $1,c
+
+The control dependencies would then extend only to the pair of cmov
+instructions and the store depending on them.  This means that a weakly
+ordered CPU would have no dependency of any sort between the load from
+"a" and the store to "c".  In short, control dependencies provide ordering
+only to the stores in the then-clause and else-clause of the "if" statement
+in question (including functions invoked by those two clauses), and not
+to code following that "if" statement.
+
+
+In summary:
+
+  (*) Control dependencies can order prior loads against later stores.
+      However, they do *not* guarantee any other sort of ordering:
+      Not prior loads against later loads, nor prior stores against
+      later anything.  If you need these other forms of ordering, use
+      smp_load_acquire(), smp_store_release(), or, in the case of prior
+      stores and later loads, smp_mb().
+
+  (*) If both legs of the "if" statement contain identical stores to
+      the same variable, then you must explicitly order those stores,
+      either by preceding both of them with smp_mb() or by using
+      smp_store_release().  Please note that it is *not* sufficient to use
+      barrier() at beginning and end of each leg of the "if" statement
+      because, as shown by the example above, optimizing compilers can
+      destroy the control dependency while respecting the letter of the
+      barrier() law.
+
+  (*) Control dependencies require at least one run-time conditional
+      between the prior load and the subsequent store, and this
+      conditional must involve the prior load.  If the compiler is able
+      to optimize the conditional away, it will have also optimized
+      away the ordering.  Careful use of READ_ONCE() and WRITE_ONCE()
+      can help to preserve the needed conditional.
+
+  (*) Control dependencies require that the compiler avoid reordering the
+      dependency into nonexistence.  Careful use of READ_ONCE() or
+      atomic{,64}_read() can help to preserve your control dependency.
+
+  (*) Control dependencies apply only to the then-clause and else-clause
+      of the "if" statement containing the control dependency, including
+      any functions that these two clauses call.  Control dependencies
+      do *not* apply to code beyond the end of that "if" statement.
+
+  (*) Control dependencies pair normally with other types of barriers.
+
+  (*) Control dependencies do *not* provide multicopy atomicity.  If you
+      need all the CPUs to agree on the ordering of a given store against
+      all other accesses, use smp_mb().
+
+  (*) Compilers do not understand control dependencies.  It is therefore
+      your job to ensure that they do not break your code.
diff --git a/tools/memory-model/Documentation/explanation.txt b/tools/memory-model/Documentation/explanation.txt
new file mode 100644
index 0000000000..6dc8b36424
--- /dev/null
+++ b/tools/memory-model/Documentation/explanation.txt
@@ -0,0 +1,2810 @@
+Explanation of the Linux-Kernel Memory Consistency Model
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+:Author: Alan Stern <stern@rowland.harvard.edu>
+:Created: October 2017
+
+.. Contents
+
+  1. INTRODUCTION
+  2. BACKGROUND
+  3. A SIMPLE EXAMPLE
+  4. A SELECTION OF MEMORY MODELS
+  5. ORDERING AND CYCLES
+  6. EVENTS
+  7. THE PROGRAM ORDER RELATION: po AND po-loc
+  8. A WARNING
+  9. DEPENDENCY RELATIONS: data, addr, and ctrl
+  10. THE READS-FROM RELATION: rf, rfi, and rfe
+  11. CACHE COHERENCE AND THE COHERENCE ORDER RELATION: co, coi, and coe
+  12. THE FROM-READS RELATION: fr, fri, and fre
+  13. AN OPERATIONAL MODEL
+  14. PROPAGATION ORDER RELATION: cumul-fence
+  15. DERIVATION OF THE LKMM FROM THE OPERATIONAL MODEL
+  16. SEQUENTIAL CONSISTENCY PER VARIABLE
+  17. ATOMIC UPDATES: rmw
+  18. THE PRESERVED PROGRAM ORDER RELATION: ppo
+  19. AND THEN THERE WAS ALPHA
+  20. THE HAPPENS-BEFORE RELATION: hb
+  21. THE PROPAGATES-BEFORE RELATION: pb
+  22. RCU RELATIONS: rcu-link, rcu-gp, rcu-rscsi, rcu-order, rcu-fence, and rb
+  23. SRCU READ-SIDE CRITICAL SECTIONS
+  24. LOCKING
+  25. PLAIN ACCESSES AND DATA RACES
+  26. ODDS AND ENDS
+
+
+
+INTRODUCTION
+------------
+
+The Linux-kernel memory consistency model (LKMM) is rather complex and
+obscure.  This is particularly evident if you read through the
+linux-kernel.bell and linux-kernel.cat files that make up the formal
+version of the model; they are extremely terse and their meanings are
+far from clear.
+
+This document describes the ideas underlying the LKMM.  It is meant
+for people who want to understand how the model was designed.  It does
+not go into the details of the code in the .bell and .cat files;
+rather, it explains in English what the code expresses symbolically.
+
+Sections 2 (BACKGROUND) through 5 (ORDERING AND CYCLES) are aimed
+toward beginners; they explain what memory consistency models are and
+the basic notions shared by all such models.  People already familiar
+with these concepts can skim or skip over them.  Sections 6 (EVENTS)
+through 12 (THE FROM_READS RELATION) describe the fundamental
+relations used in many models.  Starting in Section 13 (AN OPERATIONAL
+MODEL), the workings of the LKMM itself are covered.
+
+Warning: The code examples in this document are not written in the
+proper format for litmus tests.  They don't include a header line, the
+initializations are not enclosed in braces, the global variables are
+not passed by pointers, and they don't have an "exists" clause at the
+end.  Converting them to the right format is left as an exercise for
+the reader.
+
+
+BACKGROUND
+----------
+
+A memory consistency model (or just memory model, for short) is
+something which predicts, given a piece of computer code running on a
+particular kind of system, what values may be obtained by the code's
+load instructions.  The LKMM makes these predictions for code running
+as part of the Linux kernel.
+
+In practice, people tend to use memory models the other way around.
+That is, given a piece of code and a collection of values specified
+for the loads, the model will predict whether it is possible for the
+code to run in such a way that the loads will indeed obtain the
+specified values.  Of course, this is just another way of expressing
+the same idea.
+
+For code running on a uniprocessor system, the predictions are easy:
+Each load instruction must obtain the value written by the most recent
+store instruction accessing the same location (we ignore complicating
+factors such as DMA and mixed-size accesses.)  But on multiprocessor
+systems, with multiple CPUs making concurrent accesses to shared
+memory locations, things aren't so simple.
+
+Different architectures have differing memory models, and the Linux
+kernel supports a variety of architectures.  The LKMM has to be fairly
+permissive, in the sense that any behavior allowed by one of these
+architectures also has to be allowed by the LKMM.
+
+
+A SIMPLE EXAMPLE
+----------------
+
+Here is a simple example to illustrate the basic concepts.  Consider
+some code running as part of a device driver for an input device.  The
+driver might contain an interrupt handler which collects data from the
+device, stores it in a buffer, and sets a flag to indicate the buffer
+is full.  Running concurrently on a different CPU might be a part of
+the driver code being executed by a process in the midst of a read(2)
+system call.  This code tests the flag to see whether the buffer is
+ready, and if it is, copies the data back to userspace.  The buffer
+and the flag are memory locations shared between the two CPUs.
+
+We can abstract out the important pieces of the driver code as follows
+(the reason for using WRITE_ONCE() and READ_ONCE() instead of simple
+assignment statements is discussed later):
+
+	int buf = 0, flag = 0;
+
+	P0()
+	{
+		WRITE_ONCE(buf, 1);
+		WRITE_ONCE(flag, 1);
+	}
+
+	P1()
+	{
+		int r1;
+		int r2 = 0;
+
+		r1 = READ_ONCE(flag);
+		if (r1)
+			r2 = READ_ONCE(buf);
+	}
+
+Here the P0() function represents the interrupt handler running on one
+CPU and P1() represents the read() routine running on another.  The
+value 1 stored in buf represents input data collected from the device.
+Thus, P0 stores the data in buf and then sets flag.  Meanwhile, P1
+reads flag into the private variable r1, and if it is set, reads the
+data from buf into a second private variable r2 for copying to
+userspace.  (Presumably if flag is not set then the driver will wait a
+while and try again.)
+
+This pattern of memory accesses, where one CPU stores values to two
+shared memory locations and another CPU loads from those locations in
+the opposite order, is widely known as the "Message Passing" or MP
+pattern.  It is typical of memory access patterns in the kernel.
+
+Please note that this example code is a simplified abstraction.  Real
+buffers are usually larger than a single integer, real device drivers
+usually use sleep and wakeup mechanisms rather than polling for I/O
+completion, and real code generally doesn't bother to copy values into
+private variables before using them.  All that is beside the point;
+the idea here is simply to illustrate the overall pattern of memory
+accesses by the CPUs.
+
+A memory model will predict what values P1 might obtain for its loads
+from flag and buf, or equivalently, what values r1 and r2 might end up
+with after the code has finished running.
+
+Some predictions are trivial.  For instance, no sane memory model would
+predict that r1 = 42 or r2 = -7, because neither of those values ever
+gets stored in flag or buf.
+
+Some nontrivial predictions are nonetheless quite simple.  For
+instance, P1 might run entirely before P0 begins, in which case r1 and
+r2 will both be 0 at the end.  Or P0 might run entirely before P1
+begins, in which case r1 and r2 will both be 1.
+
+The interesting predictions concern what might happen when the two
+routines run concurrently.  One possibility is that P1 runs after P0's
+store to buf but before the store to flag.  In this case, r1 and r2
+will again both be 0.  (If P1 had been designed to read buf
+unconditionally then we would instead have r1 = 0 and r2 = 1.)
+
+However, the most interesting possibility is where r1 = 1 and r2 = 0.
+If this were to occur it would mean the driver contains a bug, because
+incorrect data would get sent to the user: 0 instead of 1.  As it
+happens, the LKMM does predict this outcome can occur, and the example
+driver code shown above is indeed buggy.
+
+
+A SELECTION OF MEMORY MODELS
+----------------------------
+
+The first widely cited memory model, and the simplest to understand,
+is Sequential Consistency.  According to this model, systems behave as
+if each CPU executed its instructions in order but with unspecified
+timing.  In other words, the instructions from the various CPUs get
+interleaved in a nondeterministic way, always according to some single
+global order that agrees with the order of the instructions in the
+program source for each CPU.  The model says that the value obtained
+by each load is simply the value written by the most recently executed
+store to the same memory location, from any CPU.
+
+For the MP example code shown above, Sequential Consistency predicts
+that the undesired result r1 = 1, r2 = 0 cannot occur.  The reasoning
+goes like this:
+
+	Since r1 = 1, P0 must store 1 to flag before P1 loads 1 from
+	it, as loads can obtain values only from earlier stores.
+
+	P1 loads from flag before loading from buf, since CPUs execute
+	their instructions in order.
+
+	P1 must load 0 from buf before P0 stores 1 to it; otherwise r2
+	would be 1 since a load obtains its value from the most recent
+	store to the same address.
+
+	P0 stores 1 to buf before storing 1 to flag, since it executes
+	its instructions in order.
+
+	Since an instruction (in this case, P0's store to flag) cannot
+	execute before itself, the specified outcome is impossible.
+
+However, real computer hardware almost never follows the Sequential
+Consistency memory model; doing so would rule out too many valuable
+performance optimizations.  On ARM and PowerPC architectures, for
+instance, the MP example code really does sometimes yield r1 = 1 and
+r2 = 0.
+
+x86 and SPARC follow yet a different memory model: TSO (Total Store
+Ordering).  This model predicts that the undesired outcome for the MP
+pattern cannot occur, but in other respects it differs from Sequential
+Consistency.  One example is the Store Buffer (SB) pattern, in which
+each CPU stores to its own shared location and then loads from the
+other CPU's location:
+
+	int x = 0, y = 0;
+
+	P0()
+	{
+		int r0;
+
+		WRITE_ONCE(x, 1);
+		r0 = READ_ONCE(y);
+	}
+
+	P1()
+	{
+		int r1;
+
+		WRITE_ONCE(y, 1);
+		r1 = READ_ONCE(x);
+	}
+
+Sequential Consistency predicts that the outcome r0 = 0, r1 = 0 is
+impossible.  (Exercise: Figure out the reasoning.)  But TSO allows
+this outcome to occur, and in fact it does sometimes occur on x86 and
+SPARC systems.
+
+The LKMM was inspired by the memory models followed by PowerPC, ARM,
+x86, Alpha, and other architectures.  However, it is different in
+detail from each of them.
+
+
+ORDERING AND CYCLES
+-------------------
+
+Memory models are all about ordering.  Often this is temporal ordering
+(i.e., the order in which certain events occur) but it doesn't have to
+be; consider for example the order of instructions in a program's
+source code.  We saw above that Sequential Consistency makes an
+important assumption that CPUs execute instructions in the same order
+as those instructions occur in the code, and there are many other
+instances of ordering playing central roles in memory models.
+
+The counterpart to ordering is a cycle.  Ordering rules out cycles:
+It's not possible to have X ordered before Y, Y ordered before Z, and
+Z ordered before X, because this would mean that X is ordered before
+itself.  The analysis of the MP example under Sequential Consistency
+involved just such an impossible cycle:
+
+	W: P0 stores 1 to flag   executes before
+	X: P1 loads 1 from flag  executes before
+	Y: P1 loads 0 from buf   executes before
+	Z: P0 stores 1 to buf    executes before
+	W: P0 stores 1 to flag.
+
+In short, if a memory model requires certain accesses to be ordered,
+and a certain outcome for the loads in a piece of code can happen only
+if those accesses would form a cycle, then the memory model predicts
+that outcome cannot occur.
+
+The LKMM is defined largely in terms of cycles, as we will see.
+
+
+EVENTS
+------
+
+The LKMM does not work directly with the C statements that make up
+kernel source code.  Instead it considers the effects of those
+statements in a more abstract form, namely, events.  The model
+includes three types of events:
+
+	Read events correspond to loads from shared memory, such as
+	calls to READ_ONCE(), smp_load_acquire(), or
+	rcu_dereference().
+
+	Write events correspond to stores to shared memory, such as
+	calls to WRITE_ONCE(), smp_store_release(), or atomic_set().
+
+	Fence events correspond to memory barriers (also known as
+	fences), such as calls to smp_rmb() or rcu_read_lock().
+
+These categories are not exclusive; a read or write event can also be
+a fence.  This happens with functions like smp_load_acquire() or
+spin_lock().  However, no single event can be both a read and a write.
+Atomic read-modify-write accesses, such as atomic_inc() or xchg(),
+correspond to a pair of events: a read followed by a write.  (The
+write event is omitted for executions where it doesn't occur, such as
+a cmpxchg() where the comparison fails.)
+
+Other parts of the code, those which do not involve interaction with
+shared memory, do not give rise to events.  Thus, arithmetic and
+logical computations, control-flow instructions, or accesses to
+private memory or CPU registers are not of central interest to the
+memory model.  They only affect the model's predictions indirectly.
+For example, an arithmetic computation might determine the value that
+gets stored to a shared memory location (or in the case of an array
+index, the address where the value gets stored), but the memory model
+is concerned only with the store itself -- its value and its address
+-- not the computation leading up to it.
+
+Events in the LKMM can be linked by various relations, which we will
+describe in the following sections.  The memory model requires certain
+of these relations to be orderings, that is, it requires them not to
+have any cycles.
+
+
+THE PROGRAM ORDER RELATION: po AND po-loc
+-----------------------------------------
+
+The most important relation between events is program order (po).  You
+can think of it as the order in which statements occur in the source
+code after branches are taken into account and loops have been
+unrolled.  A better description might be the order in which
+instructions are presented to a CPU's execution unit.  Thus, we say
+that X is po-before Y (written as "X ->po Y" in formulas) if X occurs
+before Y in the instruction stream.
+
+This is inherently a single-CPU relation; two instructions executing
+on different CPUs are never linked by po.  Also, it is by definition
+an ordering so it cannot have any cycles.
+
+po-loc is a sub-relation of po.  It links two memory accesses when the
+first comes before the second in program order and they access the
+same memory location (the "-loc" suffix).
+
+Although this may seem straightforward, there is one subtle aspect to
+program order we need to explain.  The LKMM was inspired by low-level
+architectural memory models which describe the behavior of machine
+code, and it retains their outlook to a considerable extent.  The
+read, write, and fence events used by the model are close in spirit to
+individual machine instructions.  Nevertheless, the LKMM describes
+kernel code written in C, and the mapping from C to machine code can
+be extremely complex.
+
+Optimizing compilers have great freedom in the way they translate
+source code to object code.  They are allowed to apply transformations
+that add memory accesses, eliminate accesses, combine them, split them
+into pieces, or move them around.  The use of READ_ONCE(), WRITE_ONCE(),
+or one of the other atomic or synchronization primitives prevents a
+large number of compiler optimizations.  In particular, it is guaranteed
+that the compiler will not remove such accesses from the generated code
+(unless it can prove the accesses will never be executed), it will not
+change the order in which they occur in the code (within limits imposed
+by the C standard), and it will not introduce extraneous accesses.
+
+The MP and SB examples above used READ_ONCE() and WRITE_ONCE() rather
+than ordinary memory accesses.  Thanks to this usage, we can be certain
+that in the MP example, the compiler won't reorder P0's write event to
+buf and P0's write event to flag, and similarly for the other shared
+memory accesses in the examples.
+
+Since private variables are not shared between CPUs, they can be
+accessed normally without READ_ONCE() or WRITE_ONCE().  In fact, they
+need not even be stored in normal memory at all -- in principle a
+private variable could be stored in a CPU register (hence the convention
+that these variables have names starting with the letter 'r').
+
+
+A WARNING
+---------
+
+The protections provided by READ_ONCE(), WRITE_ONCE(), and others are
+not perfect; and under some circumstances it is possible for the
+compiler to undermine the memory model.  Here is an example.  Suppose
+both branches of an "if" statement store the same value to the same
+location:
+
+	r1 = READ_ONCE(x);
+	if (r1) {
+		WRITE_ONCE(y, 2);
+		...  /* do something */
+	} else {
+		WRITE_ONCE(y, 2);
+		...  /* do something else */
+	}
+
+For this code, the LKMM predicts that the load from x will always be
+executed before either of the stores to y.  However, a compiler could
+lift the stores out of the conditional, transforming the code into
+something resembling:
+
+	r1 = READ_ONCE(x);
+	WRITE_ONCE(y, 2);
+	if (r1) {
+		...  /* do something */
+	} else {
+		...  /* do something else */
+	}
+
+Given this version of the code, the LKMM would predict that the load
+from x could be executed after the store to y.  Thus, the memory
+model's original prediction could be invalidated by the compiler.
+
+Another issue arises from the fact that in C, arguments to many
+operators and function calls can be evaluated in any order.  For
+example:
+
+	r1 = f(5) + g(6);
+
+The object code might call f(5) either before or after g(6); the
+memory model cannot assume there is a fixed program order relation
+between them.  (In fact, if the function calls are inlined then the
+compiler might even interleave their object code.)
+
+
+DEPENDENCY RELATIONS: data, addr, and ctrl
+------------------------------------------
+
+We say that two events are linked by a dependency relation when the
+execution of the second event depends in some way on a value obtained
+from memory by the first.  The first event must be a read, and the
+value it obtains must somehow affect what the second event does.
+There are three kinds of dependencies: data, address (addr), and
+control (ctrl).
+
+A read and a write event are linked by a data dependency if the value
+obtained by the read affects the value stored by the write.  As a very
+simple example:
+
+	int x, y;
+
+	r1 = READ_ONCE(x);
+	WRITE_ONCE(y, r1 + 5);
+
+The value stored by the WRITE_ONCE obviously depends on the value
+loaded by the READ_ONCE.  Such dependencies can wind through
+arbitrarily complicated computations, and a write can depend on the
+values of multiple reads.
+
+A read event and another memory access event are linked by an address
+dependency if the value obtained by the read affects the location
+accessed by the other event.  The second event can be either a read or
+a write.  Here's another simple example:
+
+	int a[20];
+	int i;
+
+	r1 = READ_ONCE(i);
+	r2 = READ_ONCE(a[r1]);
+
+Here the location accessed by the second READ_ONCE() depends on the
+index value loaded by the first.  Pointer indirection also gives rise
+to address dependencies, since the address of a location accessed
+through a pointer will depend on the value read earlier from that
+pointer.
+
+Finally, a read event X and a write event Y are linked by a control
+dependency if Y syntactically lies within an arm of an if statement and
+X affects the evaluation of the if condition via a data or address
+dependency (or similarly for a switch statement).  Simple example:
+
+	int x, y;
+
+	r1 = READ_ONCE(x);
+	if (r1)
+		WRITE_ONCE(y, 1984);
+
+Execution of the WRITE_ONCE() is controlled by a conditional expression
+which depends on the value obtained by the READ_ONCE(); hence there is
+a control dependency from the load to the store.
+
+It should be pretty obvious that events can only depend on reads that
+come earlier in program order.  Symbolically, if we have R ->data X,
+R ->addr X, or R ->ctrl X (where R is a read event), then we must also
+have R ->po X.  It wouldn't make sense for a computation to depend
+somehow on a value that doesn't get loaded from shared memory until
+later in the code!
+
+Here's a trick question: When is a dependency not a dependency?  Answer:
+When it is purely syntactic rather than semantic.  We say a dependency
+between two accesses is purely syntactic if the second access doesn't
+actually depend on the result of the first.  Here is a trivial example:
+
+	r1 = READ_ONCE(x);
+	WRITE_ONCE(y, r1 * 0);
+
+There appears to be a data dependency from the load of x to the store
+of y, since the value to be stored is computed from the value that was
+loaded.  But in fact, the value stored does not really depend on
+anything since it will always be 0.  Thus the data dependency is only
+syntactic (it appears to exist in the code) but not semantic (the
+second access will always be the same, regardless of the value of the
+first access).  Given code like this, a compiler could simply discard
+the value returned by the load from x, which would certainly destroy
+any dependency.  (The compiler is not permitted to eliminate entirely
+the load generated for a READ_ONCE() -- that's one of the nice
+properties of READ_ONCE() -- but it is allowed to ignore the load's
+value.)
+
+It's natural to object that no one in their right mind would write
+code like the above.  However, macro expansions can easily give rise
+to this sort of thing, in ways that often are not apparent to the
+programmer.
+
+Another mechanism that can lead to purely syntactic dependencies is
+related to the notion of "undefined behavior".  Certain program
+behaviors are called "undefined" in the C language specification,
+which means that when they occur there are no guarantees at all about
+the outcome.  Consider the following example:
+
+	int a[1];
+	int i;
+
+	r1 = READ_ONCE(i);
+	r2 = READ_ONCE(a[r1]);
+
+Access beyond the end or before the beginning of an array is one kind
+of undefined behavior.  Therefore the compiler doesn't have to worry
+about what will happen if r1 is nonzero, and it can assume that r1
+will always be zero regardless of the value actually loaded from i.
+(If the assumption turns out to be wrong the resulting behavior will
+be undefined anyway, so the compiler doesn't care!)  Thus the value
+from the load can be discarded, breaking the address dependency.
+
+The LKMM is unaware that purely syntactic dependencies are different
+from semantic dependencies and therefore mistakenly predicts that the
+accesses in the two examples above will be ordered.  This is another
+example of how the compiler can undermine the memory model.  Be warned.
+
+
+THE READS-FROM RELATION: rf, rfi, and rfe
+-----------------------------------------
+
+The reads-from relation (rf) links a write event to a read event when
+the value loaded by the read is the value that was stored by the
+write.  In colloquial terms, the load "reads from" the store.  We
+write W ->rf R to indicate that the load R reads from the store W.  We
+further distinguish the cases where the load and the store occur on
+the same CPU (internal reads-from, or rfi) and where they occur on
+different CPUs (external reads-from, or rfe).
+
+For our purposes, a memory location's initial value is treated as
+though it had been written there by an imaginary initial store that
+executes on a separate CPU before the main program runs.
+
+Usage of the rf relation implicitly assumes that loads will always
+read from a single store.  It doesn't apply properly in the presence
+of load-tearing, where a load obtains some of its bits from one store
+and some of them from another store.  Fortunately, use of READ_ONCE()
+and WRITE_ONCE() will prevent load-tearing; it's not possible to have:
+
+	int x = 0;
+
+	P0()
+	{
+		WRITE_ONCE(x, 0x1234);
+	}
+
+	P1()
+	{
+		int r1;
+
+		r1 = READ_ONCE(x);
+	}
+
+and end up with r1 = 0x1200 (partly from x's initial value and partly
+from the value stored by P0).
+
+On the other hand, load-tearing is unavoidable when mixed-size
+accesses are used.  Consider this example:
+
+	union {
+		u32	w;
+		u16	h[2];
+	} x;
+
+	P0()
+	{
+		WRITE_ONCE(x.h[0], 0x1234);
+		WRITE_ONCE(x.h[1], 0x5678);
+	}
+
+	P1()
+	{
+		int r1;
+
+		r1 = READ_ONCE(x.w);
+	}
+
+If r1 = 0x56781234 (little-endian!) at the end, then P1 must have read
+from both of P0's stores.  It is possible to handle mixed-size and
+unaligned accesses in a memory model, but the LKMM currently does not
+attempt to do so.  It requires all accesses to be properly aligned and
+of the location's actual size.
+
+
+CACHE COHERENCE AND THE COHERENCE ORDER RELATION: co, coi, and coe
+------------------------------------------------------------------
+
+Cache coherence is a general principle requiring that in a
+multi-processor system, the CPUs must share a consistent view of the
+memory contents.  Specifically, it requires that for each location in
+shared memory, the stores to that location must form a single global
+ordering which all the CPUs agree on (the coherence order), and this
+ordering must be consistent with the program order for accesses to
+that location.
+
+To put it another way, for any variable x, the coherence order (co) of
+the stores to x is simply the order in which the stores overwrite one
+another.  The imaginary store which establishes x's initial value
+comes first in the coherence order; the store which directly
+overwrites the initial value comes second; the store which overwrites
+that value comes third, and so on.
+
+You can think of the coherence order as being the order in which the
+stores reach x's location in memory (or if you prefer a more
+hardware-centric view, the order in which the stores get written to
+x's cache line).  We write W ->co W' if W comes before W' in the
+coherence order, that is, if the value stored by W gets overwritten,
+directly or indirectly, by the value stored by W'.
+
+Coherence order is required to be consistent with program order.  This
+requirement takes the form of four coherency rules:
+
+	Write-write coherence: If W ->po-loc W' (i.e., W comes before
+	W' in program order and they access the same location), where W
+	and W' are two stores, then W ->co W'.
+
+	Write-read coherence: If W ->po-loc R, where W is a store and R
+	is a load, then R must read from W or from some other store
+	which comes after W in the coherence order.
+
+	Read-write coherence: If R ->po-loc W, where R is a load and W
+	is a store, then the store which R reads from must come before
+	W in the coherence order.
+
+	Read-read coherence: If R ->po-loc R', where R and R' are two
+	loads, then either they read from the same store or else the
+	store read by R comes before the store read by R' in the
+	coherence order.
+
+This is sometimes referred to as sequential consistency per variable,
+because it means that the accesses to any single memory location obey
+the rules of the Sequential Consistency memory model.  (According to
+Wikipedia, sequential consistency per variable and cache coherence
+mean the same thing except that cache coherence includes an extra
+requirement that every store eventually becomes visible to every CPU.)
+
+Any reasonable memory model will include cache coherence.  Indeed, our
+expectation of cache coherence is so deeply ingrained that violations
+of its requirements look more like hardware bugs than programming
+errors:
+
+	int x;
+
+	P0()
+	{
+		WRITE_ONCE(x, 17);
+		WRITE_ONCE(x, 23);
+	}
+
+If the final value stored in x after this code ran was 17, you would
+think your computer was broken.  It would be a violation of the
+write-write coherence rule: Since the store of 23 comes later in
+program order, it must also come later in x's coherence order and
+thus must overwrite the store of 17.
+
+	int x = 0;
+
+	P0()
+	{
+		int r1;
+
+		r1 = READ_ONCE(x);
+		WRITE_ONCE(x, 666);
+	}
+
+If r1 = 666 at the end, this would violate the read-write coherence
+rule: The READ_ONCE() load comes before the WRITE_ONCE() store in
+program order, so it must not read from that store but rather from one
+coming earlier in the coherence order (in this case, x's initial
+value).
+
+	int x = 0;
+
+	P0()
+	{
+		WRITE_ONCE(x, 5);
+	}
+
+	P1()
+	{
+		int r1, r2;
+
+		r1 = READ_ONCE(x);
+		r2 = READ_ONCE(x);
+	}
+
+If r1 = 5 (reading from P0's store) and r2 = 0 (reading from the
+imaginary store which establishes x's initial value) at the end, this
+would violate the read-read coherence rule: The r1 load comes before
+the r2 load in program order, so it must not read from a store that
+comes later in the coherence order.
+
+(As a minor curiosity, if this code had used normal loads instead of
+READ_ONCE() in P1, on Itanium it sometimes could end up with r1 = 5
+and r2 = 0!  This results from parallel execution of the operations
+encoded in Itanium's Very-Long-Instruction-Word format, and it is yet
+another motivation for using READ_ONCE() when accessing shared memory
+locations.)
+
+Just like the po relation, co is inherently an ordering -- it is not
+possible for a store to directly or indirectly overwrite itself!  And
+just like with the rf relation, we distinguish between stores that
+occur on the same CPU (internal coherence order, or coi) and stores
+that occur on different CPUs (external coherence order, or coe).
+
+On the other hand, stores to different memory locations are never
+related by co, just as instructions on different CPUs are never
+related by po.  Coherence order is strictly per-location, or if you
+prefer, each location has its own independent coherence order.
+
+
+THE FROM-READS RELATION: fr, fri, and fre
+-----------------------------------------
+
+The from-reads relation (fr) can be a little difficult for people to
+grok.  It describes the situation where a load reads a value that gets
+overwritten by a store.  In other words, we have R ->fr W when the
+value that R reads is overwritten (directly or indirectly) by W, or
+equivalently, when R reads from a store which comes earlier than W in
+the coherence order.
+
+For example:
+
+	int x = 0;
+
+	P0()
+	{
+		int r1;
+
+		r1 = READ_ONCE(x);
+		WRITE_ONCE(x, 2);
+	}
+
+The value loaded from x will be 0 (assuming cache coherence!), and it
+gets overwritten by the value 2.  Thus there is an fr link from the
+READ_ONCE() to the WRITE_ONCE().  If the code contained any later
+stores to x, there would also be fr links from the READ_ONCE() to
+them.
+
+As with rf, rfi, and rfe, we subdivide the fr relation into fri (when
+the load and the store are on the same CPU) and fre (when they are on
+different CPUs).
+
+Note that the fr relation is determined entirely by the rf and co
+relations; it is not independent.  Given a read event R and a write
+event W for the same location, we will have R ->fr W if and only if
+the write which R reads from is co-before W.  In symbols,
+
+	(R ->fr W) := (there exists W' with W' ->rf R and W' ->co W).
+
+
+AN OPERATIONAL MODEL
+--------------------
+
+The LKMM is based on various operational memory models, meaning that
+the models arise from an abstract view of how a computer system
+operates.  Here are the main ideas, as incorporated into the LKMM.
+
+The system as a whole is divided into the CPUs and a memory subsystem.
+The CPUs are responsible for executing instructions (not necessarily
+in program order), and they communicate with the memory subsystem.
+For the most part, executing an instruction requires a CPU to perform
+only internal operations.  However, loads, stores, and fences involve
+more.
+
+When CPU C executes a store instruction, it tells the memory subsystem
+to store a certain value at a certain location.  The memory subsystem
+propagates the store to all the other CPUs as well as to RAM.  (As a
+special case, we say that the store propagates to its own CPU at the
+time it is executed.)  The memory subsystem also determines where the
+store falls in the location's coherence order.  In particular, it must
+arrange for the store to be co-later than (i.e., to overwrite) any
+other store to the same location which has already propagated to CPU C.
+
+When a CPU executes a load instruction R, it first checks to see
+whether there are any as-yet unexecuted store instructions, for the
+same location, that come before R in program order.  If there are, it
+uses the value of the po-latest such store as the value obtained by R,
+and we say that the store's value is forwarded to R.  Otherwise, the
+CPU asks the memory subsystem for the value to load and we say that R
+is satisfied from memory.  The memory subsystem hands back the value
+of the co-latest store to the location in question which has already
+propagated to that CPU.
+
+(In fact, the picture needs to be a little more complicated than this.
+CPUs have local caches, and propagating a store to a CPU really means
+propagating it to the CPU's local cache.  A local cache can take some
+time to process the stores that it receives, and a store can't be used
+to satisfy one of the CPU's loads until it has been processed.  On
+most architectures, the local caches process stores in
+First-In-First-Out order, and consequently the processing delay
+doesn't matter for the memory model.  But on Alpha, the local caches
+have a partitioned design that results in non-FIFO behavior.  We will
+discuss this in more detail later.)
+
+Note that load instructions may be executed speculatively and may be
+restarted under certain circumstances.  The memory model ignores these
+premature executions; we simply say that the load executes at the
+final time it is forwarded or satisfied.
+
+Executing a fence (or memory barrier) instruction doesn't require a
+CPU to do anything special other than informing the memory subsystem
+about the fence.  However, fences do constrain the way CPUs and the
+memory subsystem handle other instructions, in two respects.
+
+First, a fence forces the CPU to execute various instructions in
+program order.  Exactly which instructions are ordered depends on the
+type of fence:
+
+	Strong fences, including smp_mb() and synchronize_rcu(), force
+	the CPU to execute all po-earlier instructions before any
+	po-later instructions;
+
+	smp_rmb() forces the CPU to execute all po-earlier loads
+	before any po-later loads;
+
+	smp_wmb() forces the CPU to execute all po-earlier stores
+	before any po-later stores;
+
+	Acquire fences, such as smp_load_acquire(), force the CPU to
+	execute the load associated with the fence (e.g., the load
+	part of an smp_load_acquire()) before any po-later
+	instructions;
+
+	Release fences, such as smp_store_release(), force the CPU to
+	execute all po-earlier instructions before the store
+	associated with the fence (e.g., the store part of an
+	smp_store_release()).
+
+Second, some types of fence affect the way the memory subsystem
+propagates stores.  When a fence instruction is executed on CPU C:
+
+	For each other CPU C', smp_wmb() forces all po-earlier stores
+	on C to propagate to C' before any po-later stores do.
+
+	For each other CPU C', any store which propagates to C before
+	a release fence is executed (including all po-earlier
+	stores executed on C) is forced to propagate to C' before the
+	store associated with the release fence does.
+
+	Any store which propagates to C before a strong fence is
+	executed (including all po-earlier stores on C) is forced to
+	propagate to all other CPUs before any instructions po-after
+	the strong fence are executed on C.
+
+The propagation ordering enforced by release fences and strong fences
+affects stores from other CPUs that propagate to CPU C before the
+fence is executed, as well as stores that are executed on C before the
+fence.  We describe this property by saying that release fences and
+strong fences are A-cumulative.  By contrast, smp_wmb() fences are not
+A-cumulative; they only affect the propagation of stores that are
+executed on C before the fence (i.e., those which precede the fence in
+program order).
+
+rcu_read_lock(), rcu_read_unlock(), and synchronize_rcu() fences have
+other properties which we discuss later.
+
+
+PROPAGATION ORDER RELATION: cumul-fence
+---------------------------------------
+
+The fences which affect propagation order (i.e., strong, release, and
+smp_wmb() fences) are collectively referred to as cumul-fences, even
+though smp_wmb() isn't A-cumulative.  The cumul-fence relation is
+defined to link memory access events E and F whenever:
+
+	E and F are both stores on the same CPU and an smp_wmb() fence
+	event occurs between them in program order; or
+
+	F is a release fence and some X comes before F in program order,
+	where either X = E or else E ->rf X; or
+
+	A strong fence event occurs between some X and F in program
+	order, where either X = E or else E ->rf X.
+
+The operational model requires that whenever W and W' are both stores
+and W ->cumul-fence W', then W must propagate to any given CPU
+before W' does.  However, for different CPUs C and C', it does not
+require W to propagate to C before W' propagates to C'.
+
+
+DERIVATION OF THE LKMM FROM THE OPERATIONAL MODEL
+-------------------------------------------------
+
+The LKMM is derived from the restrictions imposed by the design
+outlined above.  These restrictions involve the necessity of
+maintaining cache coherence and the fact that a CPU can't operate on a
+value before it knows what that value is, among other things.
+
+The formal version of the LKMM is defined by six requirements, or
+axioms:
+
+	Sequential consistency per variable: This requires that the
+	system obey the four coherency rules.
+
+	Atomicity: This requires that atomic read-modify-write
+	operations really are atomic, that is, no other stores can
+	sneak into the middle of such an update.
+
+	Happens-before: This requires that certain instructions are
+	executed in a specific order.
+
+	Propagation: This requires that certain stores propagate to
+	CPUs and to RAM in a specific order.
+
+	Rcu: This requires that RCU read-side critical sections and
+	grace periods obey the rules of RCU, in particular, the
+	Grace-Period Guarantee.
+
+	Plain-coherence: This requires that plain memory accesses
+	(those not using READ_ONCE(), WRITE_ONCE(), etc.) must obey
+	the operational model's rules regarding cache coherence.
+
+The first and second are quite common; they can be found in many
+memory models (such as those for C11/C++11).  The "happens-before" and
+"propagation" axioms have analogs in other memory models as well.  The
+"rcu" and "plain-coherence" axioms are specific to the LKMM.
+
+Each of these axioms is discussed below.
+
+
+SEQUENTIAL CONSISTENCY PER VARIABLE
+-----------------------------------
+
+According to the principle of cache coherence, the stores to any fixed
+shared location in memory form a global ordering.  We can imagine
+inserting the loads from that location into this ordering, by placing
+each load between the store that it reads from and the following
+store.  This leaves the relative positions of loads that read from the
+same store unspecified; let's say they are inserted in program order,
+first for CPU 0, then CPU 1, etc.
+
+You can check that the four coherency rules imply that the rf, co, fr,
+and po-loc relations agree with this global ordering; in other words,
+whenever we have X ->rf Y or X ->co Y or X ->fr Y or X ->po-loc Y, the
+X event comes before the Y event in the global ordering.  The LKMM's
+"coherence" axiom expresses this by requiring the union of these
+relations not to have any cycles.  This means it must not be possible
+to find events
+
+	X0 -> X1 -> X2 -> ... -> Xn -> X0,
+
+where each of the links is either rf, co, fr, or po-loc.  This has to
+hold if the accesses to the fixed memory location can be ordered as
+cache coherence demands.
+
+Although it is not obvious, it can be shown that the converse is also
+true: This LKMM axiom implies that the four coherency rules are
+obeyed.
+
+
+ATOMIC UPDATES: rmw
+-------------------
+
+What does it mean to say that a read-modify-write (rmw) update, such
+as atomic_inc(&x), is atomic?  It means that the memory location (x in
+this case) does not get altered between the read and the write events
+making up the atomic operation.  In particular, if two CPUs perform
+atomic_inc(&x) concurrently, it must be guaranteed that the final
+value of x will be the initial value plus two.  We should never have
+the following sequence of events:
+
+	CPU 0 loads x obtaining 13;
+					CPU 1 loads x obtaining 13;
+	CPU 0 stores 14 to x;
+					CPU 1 stores 14 to x;
+
+where the final value of x is wrong (14 rather than 15).
+
+In this example, CPU 0's increment effectively gets lost because it
+occurs in between CPU 1's load and store.  To put it another way, the
+problem is that the position of CPU 0's store in x's coherence order
+is between the store that CPU 1 reads from and the store that CPU 1
+performs.
+
+The same analysis applies to all atomic update operations.  Therefore,
+to enforce atomicity the LKMM requires that atomic updates follow this
+rule: Whenever R and W are the read and write events composing an
+atomic read-modify-write and W' is the write event which R reads from,
+there must not be any stores coming between W' and W in the coherence
+order.  Equivalently,
+
+	(R ->rmw W) implies (there is no X with R ->fr X and X ->co W),
+
+where the rmw relation links the read and write events making up each
+atomic update.  This is what the LKMM's "atomic" axiom says.
+
+Atomic rmw updates play one more role in the LKMM: They can form "rmw
+sequences".  An rmw sequence is simply a bunch of atomic updates where
+each update reads from the previous one.  Written using events, it
+looks like this:
+
+	Z0 ->rf Y1 ->rmw Z1 ->rf ... ->rf Yn ->rmw Zn,
+
+where Z0 is some store event and n can be any number (even 0, in the
+degenerate case).  We write this relation as: Z0 ->rmw-sequence Zn.
+Note that this implies Z0 and Zn are stores to the same variable.
+
+Rmw sequences have a special property in the LKMM: They can extend the
+cumul-fence relation.  That is, if we have:
+
+	U ->cumul-fence X -> rmw-sequence Y
+
+then also U ->cumul-fence Y.  Thinking about this in terms of the
+operational model, U ->cumul-fence X says that the store U propagates
+to each CPU before the store X does.  Then the fact that X and Y are
+linked by an rmw sequence means that U also propagates to each CPU
+before Y does.  In an analogous way, rmw sequences can also extend
+the w-post-bounded relation defined below in the PLAIN ACCESSES AND
+DATA RACES section.
+
+(The notion of rmw sequences in the LKMM is similar to, but not quite
+the same as, that of release sequences in the C11 memory model.  They
+were added to the LKMM to fix an obscure bug; without them, atomic
+updates with full-barrier semantics did not always guarantee ordering
+at least as strong as atomic updates with release-barrier semantics.)
+
+
+THE PRESERVED PROGRAM ORDER RELATION: ppo
+-----------------------------------------
+
+There are many situations where a CPU is obliged to execute two
+instructions in program order.  We amalgamate them into the ppo (for
+"preserved program order") relation, which links the po-earlier
+instruction to the po-later instruction and is thus a sub-relation of
+po.
+
+The operational model already includes a description of one such
+situation: Fences are a source of ppo links.  Suppose X and Y are
+memory accesses with X ->po Y; then the CPU must execute X before Y if
+any of the following hold:
+
+	A strong (smp_mb() or synchronize_rcu()) fence occurs between
+	X and Y;
+
+	X and Y are both stores and an smp_wmb() fence occurs between
+	them;
+
+	X and Y are both loads and an smp_rmb() fence occurs between
+	them;
+
+	X is also an acquire fence, such as smp_load_acquire();
+
+	Y is also a release fence, such as smp_store_release().
+
+Another possibility, not mentioned earlier but discussed in the next
+section, is:
+
+	X and Y are both loads, X ->addr Y (i.e., there is an address
+	dependency from X to Y), and X is a READ_ONCE() or an atomic
+	access.
+
+Dependencies can also cause instructions to be executed in program
+order.  This is uncontroversial when the second instruction is a
+store; either a data, address, or control dependency from a load R to
+a store W will force the CPU to execute R before W.  This is very
+simply because the CPU cannot tell the memory subsystem about W's
+store before it knows what value should be stored (in the case of a
+data dependency), what location it should be stored into (in the case
+of an address dependency), or whether the store should actually take
+place (in the case of a control dependency).
+
+Dependencies to load instructions are more problematic.  To begin with,
+there is no such thing as a data dependency to a load.  Next, a CPU
+has no reason to respect a control dependency to a load, because it
+can always satisfy the second load speculatively before the first, and
+then ignore the result if it turns out that the second load shouldn't
+be executed after all.  And lastly, the real difficulties begin when
+we consider address dependencies to loads.
+
+To be fair about it, all Linux-supported architectures do execute
+loads in program order if there is an address dependency between them.
+After all, a CPU cannot ask the memory subsystem to load a value from
+a particular location before it knows what that location is.  However,
+the split-cache design used by Alpha can cause it to behave in a way
+that looks as if the loads were executed out of order (see the next
+section for more details).  The kernel includes a workaround for this
+problem when the loads come from READ_ONCE(), and therefore the LKMM
+includes address dependencies to loads in the ppo relation.
+
+On the other hand, dependencies can indirectly affect the ordering of
+two loads.  This happens when there is a dependency from a load to a
+store and a second, po-later load reads from that store:
+
+	R ->dep W ->rfi R',
+
+where the dep link can be either an address or a data dependency.  In
+this situation we know it is possible for the CPU to execute R' before
+W, because it can forward the value that W will store to R'.  But it
+cannot execute R' before R, because it cannot forward the value before
+it knows what that value is, or that W and R' do access the same
+location.  However, if there is merely a control dependency between R
+and W then the CPU can speculatively forward W to R' before executing
+R; if the speculation turns out to be wrong then the CPU merely has to
+restart or abandon R'.
+
+(In theory, a CPU might forward a store to a load when it runs across
+an address dependency like this:
+
+	r1 = READ_ONCE(ptr);
+	WRITE_ONCE(*r1, 17);
+	r2 = READ_ONCE(*r1);
+
+because it could tell that the store and the second load access the
+same location even before it knows what the location's address is.
+However, none of the architectures supported by the Linux kernel do
+this.)
+
+Two memory accesses of the same location must always be executed in
+program order if the second access is a store.  Thus, if we have
+
+	R ->po-loc W
+
+(the po-loc link says that R comes before W in program order and they
+access the same location), the CPU is obliged to execute W after R.
+If it executed W first then the memory subsystem would respond to R's
+read request with the value stored by W (or an even later store), in
+violation of the read-write coherence rule.  Similarly, if we had
+
+	W ->po-loc W'
+
+and the CPU executed W' before W, then the memory subsystem would put
+W' before W in the coherence order.  It would effectively cause W to
+overwrite W', in violation of the write-write coherence rule.
+(Interestingly, an early ARMv8 memory model, now obsolete, proposed
+allowing out-of-order writes like this to occur.  The model avoided
+violating the write-write coherence rule by requiring the CPU not to
+send the W write to the memory subsystem at all!)
+
+
+AND THEN THERE WAS ALPHA
+------------------------
+
+As mentioned above, the Alpha architecture is unique in that it does
+not appear to respect address dependencies to loads.  This means that
+code such as the following:
+
+	int x = 0;
+	int y = -1;
+	int *ptr = &y;
+
+	P0()
+	{
+		WRITE_ONCE(x, 1);
+		smp_wmb();
+		WRITE_ONCE(ptr, &x);
+	}
+
+	P1()
+	{
+		int *r1;
+		int r2;
+
+		r1 = ptr;
+		r2 = READ_ONCE(*r1);
+	}
+
+can malfunction on Alpha systems (notice that P1 uses an ordinary load
+to read ptr instead of READ_ONCE()).  It is quite possible that r1 = &x
+and r2 = 0 at the end, in spite of the address dependency.
+
+At first glance this doesn't seem to make sense.  We know that the
+smp_wmb() forces P0's store to x to propagate to P1 before the store
+to ptr does.  And since P1 can't execute its second load
+until it knows what location to load from, i.e., after executing its
+first load, the value x = 1 must have propagated to P1 before the
+second load executed.  So why doesn't r2 end up equal to 1?
+
+The answer lies in the Alpha's split local caches.  Although the two
+stores do reach P1's local cache in the proper order, it can happen
+that the first store is processed by a busy part of the cache while
+the second store is processed by an idle part.  As a result, the x = 1
+value may not become available for P1's CPU to read until after the
+ptr = &x value does, leading to the undesirable result above.  The
+final effect is that even though the two loads really are executed in
+program order, it appears that they aren't.
+
+This could not have happened if the local cache had processed the
+incoming stores in FIFO order.  By contrast, other architectures
+maintain at least the appearance of FIFO order.
+
+In practice, this difficulty is solved by inserting a special fence
+between P1's two loads when the kernel is compiled for the Alpha
+architecture.  In fact, as of version 4.15, the kernel automatically
+adds this fence after every READ_ONCE() and atomic load on Alpha.  The
+effect of the fence is to cause the CPU not to execute any po-later
+instructions until after the local cache has finished processing all
+the stores it has already received.  Thus, if the code was changed to:
+
+	P1()
+	{
+		int *r1;
+		int r2;
+
+		r1 = READ_ONCE(ptr);
+		r2 = READ_ONCE(*r1);
+	}
+
+then we would never get r1 = &x and r2 = 0.  By the time P1 executed
+its second load, the x = 1 store would already be fully processed by
+the local cache and available for satisfying the read request.  Thus
+we have yet another reason why shared data should always be read with
+READ_ONCE() or another synchronization primitive rather than accessed
+directly.
+
+The LKMM requires that smp_rmb(), acquire fences, and strong fences
+share this property: They do not allow the CPU to execute any po-later
+instructions (or po-later loads in the case of smp_rmb()) until all
+outstanding stores have been processed by the local cache.  In the
+case of a strong fence, the CPU first has to wait for all of its
+po-earlier stores to propagate to every other CPU in the system; then
+it has to wait for the local cache to process all the stores received
+as of that time -- not just the stores received when the strong fence
+began.
+
+And of course, none of this matters for any architecture other than
+Alpha.
+
+
+THE HAPPENS-BEFORE RELATION: hb
+-------------------------------
+
+The happens-before relation (hb) links memory accesses that have to
+execute in a certain order.  hb includes the ppo relation and two
+others, one of which is rfe.
+
+W ->rfe R implies that W and R are on different CPUs.  It also means
+that W's store must have propagated to R's CPU before R executed;
+otherwise R could not have read the value stored by W.  Therefore W
+must have executed before R, and so we have W ->hb R.
+
+The equivalent fact need not hold if W ->rfi R (i.e., W and R are on
+the same CPU).  As we have already seen, the operational model allows
+W's value to be forwarded to R in such cases, meaning that R may well
+execute before W does.
+
+It's important to understand that neither coe nor fre is included in
+hb, despite their similarities to rfe.  For example, suppose we have
+W ->coe W'.  This means that W and W' are stores to the same location,
+they execute on different CPUs, and W comes before W' in the coherence
+order (i.e., W' overwrites W).  Nevertheless, it is possible for W' to
+execute before W, because the decision as to which store overwrites
+the other is made later by the memory subsystem.  When the stores are
+nearly simultaneous, either one can come out on top.  Similarly,
+R ->fre W means that W overwrites the value which R reads, but it
+doesn't mean that W has to execute after R.  All that's necessary is
+for the memory subsystem not to propagate W to R's CPU until after R
+has executed, which is possible if W executes shortly before R.
+
+The third relation included in hb is like ppo, in that it only links
+events that are on the same CPU.  However it is more difficult to
+explain, because it arises only indirectly from the requirement of
+cache coherence.  The relation is called prop, and it links two events
+on CPU C in situations where a store from some other CPU comes after
+the first event in the coherence order and propagates to C before the
+second event executes.
+
+This is best explained with some examples.  The simplest case looks
+like this:
+
+	int x;
+
+	P0()
+	{
+		int r1;
+
+		WRITE_ONCE(x, 1);
+		r1 = READ_ONCE(x);
+	}
+
+	P1()
+	{
+		WRITE_ONCE(x, 8);
+	}
+
+If r1 = 8 at the end then P0's accesses must have executed in program
+order.  We can deduce this from the operational model; if P0's load
+had executed before its store then the value of the store would have
+been forwarded to the load, so r1 would have ended up equal to 1, not
+8.  In this case there is a prop link from P0's write event to its read
+event, because P1's store came after P0's store in x's coherence
+order, and P1's store propagated to P0 before P0's load executed.
+
+An equally simple case involves two loads of the same location that
+read from different stores:
+
+	int x = 0;
+
+	P0()
+	{
+		int r1, r2;
+
+		r1 = READ_ONCE(x);
+		r2 = READ_ONCE(x);
+	}
+
+	P1()
+	{
+		WRITE_ONCE(x, 9);
+	}
+
+If r1 = 0 and r2 = 9 at the end then P0's accesses must have executed
+in program order.  If the second load had executed before the first
+then the x = 9 store must have been propagated to P0 before the first
+load executed, and so r1 would have been 9 rather than 0.  In this
+case there is a prop link from P0's first read event to its second,
+because P1's store overwrote the value read by P0's first load, and
+P1's store propagated to P0 before P0's second load executed.
+
+Less trivial examples of prop all involve fences.  Unlike the simple
+examples above, they can require that some instructions are executed
+out of program order.  This next one should look familiar:
+
+	int buf = 0, flag = 0;
+
+	P0()
+	{
+		WRITE_ONCE(buf, 1);
+		smp_wmb();
+		WRITE_ONCE(flag, 1);
+	}
+
+	P1()
+	{
+		int r1;
+		int r2;
+
+		r1 = READ_ONCE(flag);
+		r2 = READ_ONCE(buf);
+	}
+
+This is the MP pattern again, with an smp_wmb() fence between the two
+stores.  If r1 = 1 and r2 = 0 at the end then there is a prop link
+from P1's second load to its first (backwards!).  The reason is
+similar to the previous examples: The value P1 loads from buf gets
+overwritten by P0's store to buf, the fence guarantees that the store
+to buf will propagate to P1 before the store to flag does, and the
+store to flag propagates to P1 before P1 reads flag.
+
+The prop link says that in order to obtain the r1 = 1, r2 = 0 result,
+P1 must execute its second load before the first.  Indeed, if the load
+from flag were executed first, then the buf = 1 store would already
+have propagated to P1 by the time P1's load from buf executed, so r2
+would have been 1 at the end, not 0.  (The reasoning holds even for
+Alpha, although the details are more complicated and we will not go
+into them.)
+
+But what if we put an smp_rmb() fence between P1's loads?  The fence
+would force the two loads to be executed in program order, and it
+would generate a cycle in the hb relation: The fence would create a ppo
+link (hence an hb link) from the first load to the second, and the
+prop relation would give an hb link from the second load to the first.
+Since an instruction can't execute before itself, we are forced to
+conclude that if an smp_rmb() fence is added, the r1 = 1, r2 = 0
+outcome is impossible -- as it should be.
+
+The formal definition of the prop relation involves a coe or fre link,
+followed by an arbitrary number of cumul-fence links, ending with an
+rfe link.  You can concoct more exotic examples, containing more than
+one fence, although this quickly leads to diminishing returns in terms
+of complexity.  For instance, here's an example containing a coe link
+followed by two cumul-fences and an rfe link, utilizing the fact that
+release fences are A-cumulative:
+
+	int x, y, z;
+
+	P0()
+	{
+		int r0;
+
+		WRITE_ONCE(x, 1);
+		r0 = READ_ONCE(z);
+	}
+
+	P1()
+	{
+		WRITE_ONCE(x, 2);
+		smp_wmb();
+		WRITE_ONCE(y, 1);
+	}
+
+	P2()
+	{
+		int r2;
+
+		r2 = READ_ONCE(y);
+		smp_store_release(&z, 1);
+	}
+
+If x = 2, r0 = 1, and r2 = 1 after this code runs then there is a prop
+link from P0's store to its load.  This is because P0's store gets
+overwritten by P1's store since x = 2 at the end (a coe link), the
+smp_wmb() ensures that P1's store to x propagates to P2 before the
+store to y does (the first cumul-fence), the store to y propagates to P2
+before P2's load and store execute, P2's smp_store_release()
+guarantees that the stores to x and y both propagate to P0 before the
+store to z does (the second cumul-fence), and P0's load executes after the
+store to z has propagated to P0 (an rfe link).
+
+In summary, the fact that the hb relation links memory access events
+in the order they execute means that it must not have cycles.  This
+requirement is the content of the LKMM's "happens-before" axiom.
+
+The LKMM defines yet another relation connected to times of
+instruction execution, but it is not included in hb.  It relies on the
+particular properties of strong fences, which we cover in the next
+section.
+
+
+THE PROPAGATES-BEFORE RELATION: pb
+----------------------------------
+
+The propagates-before (pb) relation capitalizes on the special
+features of strong fences.  It links two events E and F whenever some
+store is coherence-later than E and propagates to every CPU and to RAM
+before F executes.  The formal definition requires that E be linked to
+F via a coe or fre link, an arbitrary number of cumul-fences, an
+optional rfe link, a strong fence, and an arbitrary number of hb
+links.  Let's see how this definition works out.
+
+Consider first the case where E is a store (implying that the sequence
+of links begins with coe).  Then there are events W, X, Y, and Z such
+that:
+
+	E ->coe W ->cumul-fence* X ->rfe? Y ->strong-fence Z ->hb* F,
+
+where the * suffix indicates an arbitrary number of links of the
+specified type, and the ? suffix indicates the link is optional (Y may
+be equal to X).  Because of the cumul-fence links, we know that W will
+propagate to Y's CPU before X does, hence before Y executes and hence
+before the strong fence executes.  Because this fence is strong, we
+know that W will propagate to every CPU and to RAM before Z executes.
+And because of the hb links, we know that Z will execute before F.
+Thus W, which comes later than E in the coherence order, will
+propagate to every CPU and to RAM before F executes.
+
+The case where E is a load is exactly the same, except that the first
+link in the sequence is fre instead of coe.
+
+The existence of a pb link from E to F implies that E must execute
+before F.  To see why, suppose that F executed first.  Then W would
+have propagated to E's CPU before E executed.  If E was a store, the
+memory subsystem would then be forced to make E come after W in the
+coherence order, contradicting the fact that E ->coe W.  If E was a
+load, the memory subsystem would then be forced to satisfy E's read
+request with the value stored by W or an even later store,
+contradicting the fact that E ->fre W.
+
+A good example illustrating how pb works is the SB pattern with strong
+fences:
+
+	int x = 0, y = 0;
+
+	P0()
+	{
+		int r0;
+
+		WRITE_ONCE(x, 1);
+		smp_mb();
+		r0 = READ_ONCE(y);
+	}
+
+	P1()
+	{
+		int r1;
+
+		WRITE_ONCE(y, 1);
+		smp_mb();
+		r1 = READ_ONCE(x);
+	}
+
+If r0 = 0 at the end then there is a pb link from P0's load to P1's
+load: an fre link from P0's load to P1's store (which overwrites the
+value read by P0), and a strong fence between P1's store and its load.
+In this example, the sequences of cumul-fence and hb links are empty.
+Note that this pb link is not included in hb as an instance of prop,
+because it does not start and end on the same CPU.
+
+Similarly, if r1 = 0 at the end then there is a pb link from P1's load
+to P0's.  This means that if both r1 and r2 were 0 there would be a
+cycle in pb, which is not possible since an instruction cannot execute
+before itself.  Thus, adding smp_mb() fences to the SB pattern
+prevents the r0 = 0, r1 = 0 outcome.
+
+In summary, the fact that the pb relation links events in the order
+they execute means that it cannot have cycles.  This requirement is
+the content of the LKMM's "propagation" axiom.
+
+
+RCU RELATIONS: rcu-link, rcu-gp, rcu-rscsi, rcu-order, rcu-fence, and rb
+------------------------------------------------------------------------
+
+RCU (Read-Copy-Update) is a powerful synchronization mechanism.  It
+rests on two concepts: grace periods and read-side critical sections.
+
+A grace period is the span of time occupied by a call to
+synchronize_rcu().  A read-side critical section (or just critical
+section, for short) is a region of code delimited by rcu_read_lock()
+at the start and rcu_read_unlock() at the end.  Critical sections can
+be nested, although we won't make use of this fact.
+
+As far as memory models are concerned, RCU's main feature is its
+Grace-Period Guarantee, which states that a critical section can never
+span a full grace period.  In more detail, the Guarantee says:
+
+	For any critical section C and any grace period G, at least
+	one of the following statements must hold:
+
+(1)	C ends before G does, and in addition, every store that
+	propagates to C's CPU before the end of C must propagate to
+	every CPU before G ends.
+
+(2)	G starts before C does, and in addition, every store that
+	propagates to G's CPU before the start of G must propagate
+	to every CPU before C starts.
+
+In particular, it is not possible for a critical section to both start
+before and end after a grace period.
+
+Here is a simple example of RCU in action:
+
+	int x, y;
+
+	P0()
+	{
+		rcu_read_lock();
+		WRITE_ONCE(x, 1);
+		WRITE_ONCE(y, 1);
+		rcu_read_unlock();
+	}
+
+	P1()
+	{
+		int r1, r2;
+
+		r1 = READ_ONCE(x);
+		synchronize_rcu();
+		r2 = READ_ONCE(y);
+	}
+
+The Grace Period Guarantee tells us that when this code runs, it will
+never end with r1 = 1 and r2 = 0.  The reasoning is as follows.  r1 = 1
+means that P0's store to x propagated to P1 before P1 called
+synchronize_rcu(), so P0's critical section must have started before
+P1's grace period, contrary to part (2) of the Guarantee.  On the
+other hand, r2 = 0 means that P0's store to y, which occurs before the
+end of the critical section, did not propagate to P1 before the end of
+the grace period, contrary to part (1).  Together the results violate
+the Guarantee.
+
+In the kernel's implementations of RCU, the requirements for stores
+to propagate to every CPU are fulfilled by placing strong fences at
+suitable places in the RCU-related code.  Thus, if a critical section
+starts before a grace period does then the critical section's CPU will
+execute an smp_mb() fence after the end of the critical section and
+some time before the grace period's synchronize_rcu() call returns.
+And if a critical section ends after a grace period does then the
+synchronize_rcu() routine will execute an smp_mb() fence at its start
+and some time before the critical section's opening rcu_read_lock()
+executes.
+
+What exactly do we mean by saying that a critical section "starts
+before" or "ends after" a grace period?  Some aspects of the meaning
+are pretty obvious, as in the example above, but the details aren't
+entirely clear.  The LKMM formalizes this notion by means of the
+rcu-link relation.  rcu-link encompasses a very general notion of
+"before": If E and F are RCU fence events (i.e., rcu_read_lock(),
+rcu_read_unlock(), or synchronize_rcu()) then among other things,
+E ->rcu-link F includes cases where E is po-before some memory-access
+event X, F is po-after some memory-access event Y, and we have any of
+X ->rfe Y, X ->co Y, or X ->fr Y.
+
+The formal definition of the rcu-link relation is more than a little
+obscure, and we won't give it here.  It is closely related to the pb
+relation, and the details don't matter unless you want to comb through
+a somewhat lengthy formal proof.  Pretty much all you need to know
+about rcu-link is the information in the preceding paragraph.
+
+The LKMM also defines the rcu-gp and rcu-rscsi relations.  They bring
+grace periods and read-side critical sections into the picture, in the
+following way:
+
+	E ->rcu-gp F means that E and F are in fact the same event,
+	and that event is a synchronize_rcu() fence (i.e., a grace
+	period).
+
+	E ->rcu-rscsi F means that E and F are the rcu_read_unlock()
+	and rcu_read_lock() fence events delimiting some read-side
+	critical section.  (The 'i' at the end of the name emphasizes
+	that this relation is "inverted": It links the end of the
+	critical section to the start.)
+
+If we think of the rcu-link relation as standing for an extended
+"before", then X ->rcu-gp Y ->rcu-link Z roughly says that X is a
+grace period which ends before Z begins.  (In fact it covers more than
+this, because it also includes cases where some store propagates to
+Z's CPU before Z begins but doesn't propagate to some other CPU until
+after X ends.)  Similarly, X ->rcu-rscsi Y ->rcu-link Z says that X is
+the end of a critical section which starts before Z begins.
+
+The LKMM goes on to define the rcu-order relation as a sequence of
+rcu-gp and rcu-rscsi links separated by rcu-link links, in which the
+number of rcu-gp links is >= the number of rcu-rscsi links.  For
+example:
+
+	X ->rcu-gp Y ->rcu-link Z ->rcu-rscsi T ->rcu-link U ->rcu-gp V
+
+would imply that X ->rcu-order V, because this sequence contains two
+rcu-gp links and one rcu-rscsi link.  (It also implies that
+X ->rcu-order T and Z ->rcu-order V.)  On the other hand:
+
+	X ->rcu-rscsi Y ->rcu-link Z ->rcu-rscsi T ->rcu-link U ->rcu-gp V
+
+does not imply X ->rcu-order V, because the sequence contains only
+one rcu-gp link but two rcu-rscsi links.
+
+The rcu-order relation is important because the Grace Period Guarantee
+means that rcu-order links act kind of like strong fences.  In
+particular, E ->rcu-order F implies not only that E begins before F
+ends, but also that any write po-before E will propagate to every CPU
+before any instruction po-after F can execute.  (However, it does not
+imply that E must execute before F; in fact, each synchronize_rcu()
+fence event is linked to itself by rcu-order as a degenerate case.)
+
+To prove this in full generality requires some intellectual effort.
+We'll consider just a very simple case:
+
+	G ->rcu-gp W ->rcu-link Z ->rcu-rscsi F.
+
+This formula means that G and W are the same event (a grace period),
+and there are events X, Y and a read-side critical section C such that:
+
+	1. G = W is po-before or equal to X;
+
+	2. X comes "before" Y in some sense (including rfe, co and fr);
+
+	3. Y is po-before Z;
+
+	4. Z is the rcu_read_unlock() event marking the end of C;
+
+	5. F is the rcu_read_lock() event marking the start of C.
+
+From 1 - 4 we deduce that the grace period G ends before the critical
+section C.  Then part (2) of the Grace Period Guarantee says not only
+that G starts before C does, but also that any write which executes on
+G's CPU before G starts must propagate to every CPU before C starts.
+In particular, the write propagates to every CPU before F finishes
+executing and hence before any instruction po-after F can execute.
+This sort of reasoning can be extended to handle all the situations
+covered by rcu-order.
+
+The rcu-fence relation is a simple extension of rcu-order.  While
+rcu-order only links certain fence events (calls to synchronize_rcu(),
+rcu_read_lock(), or rcu_read_unlock()), rcu-fence links any events
+that are separated by an rcu-order link.  This is analogous to the way
+the strong-fence relation links events that are separated by an
+smp_mb() fence event (as mentioned above, rcu-order links act kind of
+like strong fences).  Written symbolically, X ->rcu-fence Y means
+there are fence events E and F such that:
+
+	X ->po E ->rcu-order F ->po Y.
+
+From the discussion above, we see this implies not only that X
+executes before Y, but also (if X is a store) that X propagates to
+every CPU before Y executes.  Thus rcu-fence is sort of a
+"super-strong" fence: Unlike the original strong fences (smp_mb() and
+synchronize_rcu()), rcu-fence is able to link events on different
+CPUs.  (Perhaps this fact should lead us to say that rcu-fence isn't
+really a fence at all!)
+
+Finally, the LKMM defines the RCU-before (rb) relation in terms of
+rcu-fence.  This is done in essentially the same way as the pb
+relation was defined in terms of strong-fence.  We will omit the
+details; the end result is that E ->rb F implies E must execute
+before F, just as E ->pb F does (and for much the same reasons).
+
+Putting this all together, the LKMM expresses the Grace Period
+Guarantee by requiring that the rb relation does not contain a cycle.
+Equivalently, this "rcu" axiom requires that there are no events E
+and F with E ->rcu-link F ->rcu-order E.  Or to put it a third way,
+the axiom requires that there are no cycles consisting of rcu-gp and
+rcu-rscsi alternating with rcu-link, where the number of rcu-gp links
+is >= the number of rcu-rscsi links.
+
+Justifying the axiom isn't easy, but it is in fact a valid
+formalization of the Grace Period Guarantee.  We won't attempt to go
+through the detailed argument, but the following analysis gives a
+taste of what is involved.  Suppose both parts of the Guarantee are
+violated: A critical section starts before a grace period, and some
+store propagates to the critical section's CPU before the end of the
+critical section but doesn't propagate to some other CPU until after
+the end of the grace period.
+
+Putting symbols to these ideas, let L and U be the rcu_read_lock() and
+rcu_read_unlock() fence events delimiting the critical section in
+question, and let S be the synchronize_rcu() fence event for the grace
+period.  Saying that the critical section starts before S means there
+are events Q and R where Q is po-after L (which marks the start of the
+critical section), Q is "before" R in the sense used by the rcu-link
+relation, and R is po-before the grace period S.  Thus we have:
+
+	L ->rcu-link S.
+
+Let W be the store mentioned above, let Y come before the end of the
+critical section and witness that W propagates to the critical
+section's CPU by reading from W, and let Z on some arbitrary CPU be a
+witness that W has not propagated to that CPU, where Z happens after
+some event X which is po-after S.  Symbolically, this amounts to:
+
+	S ->po X ->hb* Z ->fr W ->rf Y ->po U.
+
+The fr link from Z to W indicates that W has not propagated to Z's CPU
+at the time that Z executes.  From this, it can be shown (see the
+discussion of the rcu-link relation earlier) that S and U are related
+by rcu-link:
+
+	S ->rcu-link U.
+
+Since S is a grace period we have S ->rcu-gp S, and since L and U are
+the start and end of the critical section C we have U ->rcu-rscsi L.
+From this we obtain:
+
+	S ->rcu-gp S ->rcu-link U ->rcu-rscsi L ->rcu-link S,
+
+a forbidden cycle.  Thus the "rcu" axiom rules out this violation of
+the Grace Period Guarantee.
+
+For something a little more down-to-earth, let's see how the axiom
+works out in practice.  Consider the RCU code example from above, this
+time with statement labels added:
+
+	int x, y;
+
+	P0()
+	{
+		L: rcu_read_lock();
+		X: WRITE_ONCE(x, 1);
+		Y: WRITE_ONCE(y, 1);
+		U: rcu_read_unlock();
+	}
+
+	P1()
+	{
+		int r1, r2;
+
+		Z: r1 = READ_ONCE(x);
+		S: synchronize_rcu();
+		W: r2 = READ_ONCE(y);
+	}
+
+
+If r2 = 0 at the end then P0's store at Y overwrites the value that
+P1's load at W reads from, so we have W ->fre Y.  Since S ->po W and
+also Y ->po U, we get S ->rcu-link U.  In addition, S ->rcu-gp S
+because S is a grace period.
+
+If r1 = 1 at the end then P1's load at Z reads from P0's store at X,
+so we have X ->rfe Z.  Together with L ->po X and Z ->po S, this
+yields L ->rcu-link S.  And since L and U are the start and end of a
+critical section, we have U ->rcu-rscsi L.
+
+Then U ->rcu-rscsi L ->rcu-link S ->rcu-gp S ->rcu-link U is a
+forbidden cycle, violating the "rcu" axiom.  Hence the outcome is not
+allowed by the LKMM, as we would expect.
+
+For contrast, let's see what can happen in a more complicated example:
+
+	int x, y, z;
+
+	P0()
+	{
+		int r0;
+
+		L0: rcu_read_lock();
+		    r0 = READ_ONCE(x);
+		    WRITE_ONCE(y, 1);
+		U0: rcu_read_unlock();
+	}
+
+	P1()
+	{
+		int r1;
+
+		    r1 = READ_ONCE(y);
+		S1: synchronize_rcu();
+		    WRITE_ONCE(z, 1);
+	}
+
+	P2()
+	{
+		int r2;
+
+		L2: rcu_read_lock();
+		    r2 = READ_ONCE(z);
+		    WRITE_ONCE(x, 1);
+		U2: rcu_read_unlock();
+	}
+
+If r0 = r1 = r2 = 1 at the end, then similar reasoning to before shows
+that U0 ->rcu-rscsi L0 ->rcu-link S1 ->rcu-gp S1 ->rcu-link U2 ->rcu-rscsi
+L2 ->rcu-link U0.  However this cycle is not forbidden, because the
+sequence of relations contains fewer instances of rcu-gp (one) than of
+rcu-rscsi (two).  Consequently the outcome is allowed by the LKMM.
+The following instruction timing diagram shows how it might actually
+occur:
+
+P0			P1			P2
+--------------------	--------------------	--------------------
+rcu_read_lock()
+WRITE_ONCE(y, 1)
+			r1 = READ_ONCE(y)
+			synchronize_rcu() starts
+			.			rcu_read_lock()
+			.			WRITE_ONCE(x, 1)
+r0 = READ_ONCE(x)	.
+rcu_read_unlock()	.
+			synchronize_rcu() ends
+			WRITE_ONCE(z, 1)
+						r2 = READ_ONCE(z)
+						rcu_read_unlock()
+
+This requires P0 and P2 to execute their loads and stores out of
+program order, but of course they are allowed to do so.  And as you
+can see, the Grace Period Guarantee is not violated: The critical
+section in P0 both starts before P1's grace period does and ends
+before it does, and the critical section in P2 both starts after P1's
+grace period does and ends after it does.
+
+The LKMM supports SRCU (Sleepable Read-Copy-Update) in addition to
+normal RCU.  The ideas involved are much the same as above, with new
+relations srcu-gp and srcu-rscsi added to represent SRCU grace periods
+and read-side critical sections.  However, there are some significant
+differences between RCU read-side critical sections and their SRCU
+counterparts, as described in the next section.
+
+
+SRCU READ-SIDE CRITICAL SECTIONS
+--------------------------------
+
+The LKMM uses the srcu-rscsi relation to model SRCU read-side critical
+sections.  They differ from RCU read-side critical sections in the
+following respects:
+
+1.	Unlike the analogous RCU primitives, synchronize_srcu(),
+	srcu_read_lock(), and srcu_read_unlock() take a pointer to a
+	struct srcu_struct as an argument.  This structure is called
+	an SRCU domain, and calls linked by srcu-rscsi must have the
+	same domain.  Read-side critical sections and grace periods
+	associated with different domains are independent of one
+	another; the SRCU version of the RCU Guarantee applies only
+	to pairs of critical sections and grace periods having the
+	same domain.
+
+2.	srcu_read_lock() returns a value, called the index, which must
+	be passed to the matching srcu_read_unlock() call.  Unlike
+	rcu_read_lock() and rcu_read_unlock(), an srcu_read_lock()
+	call does not always have to match the next unpaired
+	srcu_read_unlock().  In fact, it is possible for two SRCU
+	read-side critical sections to overlap partially, as in the
+	following example (where s is an srcu_struct and idx1 and idx2
+	are integer variables):
+
+		idx1 = srcu_read_lock(&s);	// Start of first RSCS
+		idx2 = srcu_read_lock(&s);	// Start of second RSCS
+		srcu_read_unlock(&s, idx1);	// End of first RSCS
+		srcu_read_unlock(&s, idx2);	// End of second RSCS
+
+	The matching is determined entirely by the domain pointer and
+	index value.  By contrast, if the calls had been
+	rcu_read_lock() and rcu_read_unlock() then they would have
+	created two nested (fully overlapping) read-side critical
+	sections: an inner one and an outer one.
+
+3.	The srcu_down_read() and srcu_up_read() primitives work
+	exactly like srcu_read_lock() and srcu_read_unlock(), except
+	that matching calls don't have to execute on the same CPU.
+	(The names are meant to be suggestive of operations on
+	semaphores.)  Since the matching is determined by the domain
+	pointer and index value, these primitives make it possible for
+	an SRCU read-side critical section to start on one CPU and end
+	on another, so to speak.
+
+In order to account for these properties of SRCU, the LKMM models
+srcu_read_lock() as a special type of load event (which is
+appropriate, since it takes a memory location as argument and returns
+a value, just as a load does) and srcu_read_unlock() as a special type
+of store event (again appropriate, since it takes as arguments a
+memory location and a value).  These loads and stores are annotated as
+belonging to the "srcu-lock" and "srcu-unlock" event classes
+respectively.
+
+This approach allows the LKMM to tell whether two events are
+associated with the same SRCU domain, simply by checking whether they
+access the same memory location (i.e., they are linked by the loc
+relation).  It also gives a way to tell which unlock matches a
+particular lock, by checking for the presence of a data dependency
+from the load (srcu-lock) to the store (srcu-unlock).  For example,
+given the situation outlined earlier (with statement labels added):
+
+	A: idx1 = srcu_read_lock(&s);
+	B: idx2 = srcu_read_lock(&s);
+	C: srcu_read_unlock(&s, idx1);
+	D: srcu_read_unlock(&s, idx2);
+
+the LKMM will treat A and B as loads from s yielding values saved in
+idx1 and idx2 respectively.  Similarly, it will treat C and D as
+though they stored the values from idx1 and idx2 in s.  The end result
+is much as if we had written:
+
+	A: idx1 = READ_ONCE(s);
+	B: idx2 = READ_ONCE(s);
+	C: WRITE_ONCE(s, idx1);
+	D: WRITE_ONCE(s, idx2);
+
+except for the presence of the special srcu-lock and srcu-unlock
+annotations.  You can see at once that we have A ->data C and
+B ->data D.  These dependencies tell the LKMM that C is the
+srcu-unlock event matching srcu-lock event A, and D is the
+srcu-unlock event matching srcu-lock event B.
+
+This approach is admittedly a hack, and it has the potential to lead
+to problems.  For example, in:
+
+	idx1 = srcu_read_lock(&s);
+	srcu_read_unlock(&s, idx1);
+	idx2 = srcu_read_lock(&s);
+	srcu_read_unlock(&s, idx2);
+
+the LKMM will believe that idx2 must have the same value as idx1,
+since it reads from the immediately preceding store of idx1 in s.
+Fortunately this won't matter, assuming that litmus tests never do
+anything with SRCU index values other than pass them to
+srcu_read_unlock() or srcu_up_read() calls.
+
+However, sometimes it is necessary to store an index value in a
+shared variable temporarily.  In fact, this is the only way for
+srcu_down_read() to pass the index it gets to an srcu_up_read() call
+on a different CPU.  In more detail, we might have soething like:
+
+	struct srcu_struct s;
+	int x;
+
+	P0()
+	{
+		int r0;
+
+		A: r0 = srcu_down_read(&s);
+		B: WRITE_ONCE(x, r0);
+	}
+
+	P1()
+	{
+		int r1;
+
+		C: r1 = READ_ONCE(x);
+		D: srcu_up_read(&s, r1);
+	}
+
+Assuming that P1 executes after P0 and does read the index value
+stored in x, we can write this (using brackets to represent event
+annotations) as:
+
+	A[srcu-lock] ->data B[once] ->rf C[once] ->data D[srcu-unlock].
+
+The LKMM defines a carry-srcu-data relation to express this pattern;
+it permits an arbitrarily long sequence of
+
+	data ; rf
+
+pairs (that is, a data link followed by an rf link) to occur between
+an srcu-lock event and the final data dependency leading to the
+matching srcu-unlock event.  carry-srcu-data is complicated by the
+need to ensure that none of the intermediate store events in this
+sequence are instances of srcu-unlock.  This is necessary because in a
+pattern like the one above:
+
+	A: idx1 = srcu_read_lock(&s);
+	B: srcu_read_unlock(&s, idx1);
+	C: idx2 = srcu_read_lock(&s);
+	D: srcu_read_unlock(&s, idx2);
+
+the LKMM treats B as a store to the variable s and C as a load from
+that variable, creating an undesirable rf link from B to C:
+
+	A ->data B ->rf C ->data D.
+
+This would cause carry-srcu-data to mistakenly extend a data
+dependency from A to D, giving the impression that D was the
+srcu-unlock event matching A's srcu-lock.  To avoid such problems,
+carry-srcu-data does not accept sequences in which the ends of any of
+the intermediate ->data links (B above) is an srcu-unlock event.
+
+
+LOCKING
+-------
+
+The LKMM includes locking.  In fact, there is special code for locking
+in the formal model, added in order to make tools run faster.
+However, this special code is intended to be more or less equivalent
+to concepts we have already covered.  A spinlock_t variable is treated
+the same as an int, and spin_lock(&s) is treated almost the same as:
+
+	while (cmpxchg_acquire(&s, 0, 1) != 0)
+		cpu_relax();
+
+This waits until s is equal to 0 and then atomically sets it to 1,
+and the read part of the cmpxchg operation acts as an acquire fence.
+An alternate way to express the same thing would be:
+
+	r = xchg_acquire(&s, 1);
+
+along with a requirement that at the end, r = 0.  Similarly,
+spin_trylock(&s) is treated almost the same as:
+
+	return !cmpxchg_acquire(&s, 0, 1);
+
+which atomically sets s to 1 if it is currently equal to 0 and returns
+true if it succeeds (the read part of the cmpxchg operation acts as an
+acquire fence only if the operation is successful).  spin_unlock(&s)
+is treated almost the same as:
+
+	smp_store_release(&s, 0);
+
+The "almost" qualifiers above need some explanation.  In the LKMM, the
+store-release in a spin_unlock() and the load-acquire which forms the
+first half of the atomic rmw update in a spin_lock() or a successful
+spin_trylock() -- we can call these things lock-releases and
+lock-acquires -- have two properties beyond those of ordinary releases
+and acquires.
+
+First, when a lock-acquire reads from or is po-after a lock-release,
+the LKMM requires that every instruction po-before the lock-release
+must execute before any instruction po-after the lock-acquire.  This
+would naturally hold if the release and acquire operations were on
+different CPUs and accessed the same lock variable, but the LKMM says
+it also holds when they are on the same CPU, even if they access
+different lock variables.  For example:
+
+	int x, y;
+	spinlock_t s, t;
+
+	P0()
+	{
+		int r1, r2;
+
+		spin_lock(&s);
+		r1 = READ_ONCE(x);
+		spin_unlock(&s);
+		spin_lock(&t);
+		r2 = READ_ONCE(y);
+		spin_unlock(&t);
+	}
+
+	P1()
+	{
+		WRITE_ONCE(y, 1);
+		smp_wmb();
+		WRITE_ONCE(x, 1);
+	}
+
+Here the second spin_lock() is po-after the first spin_unlock(), and
+therefore the load of x must execute before the load of y, even though
+the two locking operations use different locks.  Thus we cannot have
+r1 = 1 and r2 = 0 at the end (this is an instance of the MP pattern).
+
+This requirement does not apply to ordinary release and acquire
+fences, only to lock-related operations.  For instance, suppose P0()
+in the example had been written as:
+
+	P0()
+	{
+		int r1, r2, r3;
+
+		r1 = READ_ONCE(x);
+		smp_store_release(&s, 1);
+		r3 = smp_load_acquire(&s);
+		r2 = READ_ONCE(y);
+	}
+
+Then the CPU would be allowed to forward the s = 1 value from the
+smp_store_release() to the smp_load_acquire(), executing the
+instructions in the following order:
+
+		r3 = smp_load_acquire(&s);	// Obtains r3 = 1
+		r2 = READ_ONCE(y);
+		r1 = READ_ONCE(x);
+		smp_store_release(&s, 1);	// Value is forwarded
+
+and thus it could load y before x, obtaining r2 = 0 and r1 = 1.
+
+Second, when a lock-acquire reads from or is po-after a lock-release,
+and some other stores W and W' occur po-before the lock-release and
+po-after the lock-acquire respectively, the LKMM requires that W must
+propagate to each CPU before W' does.  For example, consider:
+
+	int x, y;
+	spinlock_t s;
+
+	P0()
+	{
+		spin_lock(&s);
+		WRITE_ONCE(x, 1);
+		spin_unlock(&s);
+	}
+
+	P1()
+	{
+		int r1;
+
+		spin_lock(&s);
+		r1 = READ_ONCE(x);
+		WRITE_ONCE(y, 1);
+		spin_unlock(&s);
+	}
+
+	P2()
+	{
+		int r2, r3;
+
+		r2 = READ_ONCE(y);
+		smp_rmb();
+		r3 = READ_ONCE(x);
+	}
+
+If r1 = 1 at the end then the spin_lock() in P1 must have read from
+the spin_unlock() in P0.  Hence the store to x must propagate to P2
+before the store to y does, so we cannot have r2 = 1 and r3 = 0.  But
+if P1 had used a lock variable different from s, the writes could have
+propagated in either order.  (On the other hand, if the code in P0 and
+P1 had all executed on a single CPU, as in the example before this
+one, then the writes would have propagated in order even if the two
+critical sections used different lock variables.)
+
+These two special requirements for lock-release and lock-acquire do
+not arise from the operational model.  Nevertheless, kernel developers
+have come to expect and rely on them because they do hold on all
+architectures supported by the Linux kernel, albeit for various
+differing reasons.
+
+
+PLAIN ACCESSES AND DATA RACES
+-----------------------------
+
+In the LKMM, memory accesses such as READ_ONCE(x), atomic_inc(&y),
+smp_load_acquire(&z), and so on are collectively referred to as
+"marked" accesses, because they are all annotated with special
+operations of one kind or another.  Ordinary C-language memory
+accesses such as x or y = 0 are simply called "plain" accesses.
+
+Early versions of the LKMM had nothing to say about plain accesses.
+The C standard allows compilers to assume that the variables affected
+by plain accesses are not concurrently read or written by any other
+threads or CPUs.  This leaves compilers free to implement all manner
+of transformations or optimizations of code containing plain accesses,
+making such code very difficult for a memory model to handle.
+
+Here is just one example of a possible pitfall:
+
+	int a = 6;
+	int *x = &a;
+
+	P0()
+	{
+		int *r1;
+		int r2 = 0;
+
+		r1 = x;
+		if (r1 != NULL)
+			r2 = READ_ONCE(*r1);
+	}
+
+	P1()
+	{
+		WRITE_ONCE(x, NULL);
+	}
+
+On the face of it, one would expect that when this code runs, the only
+possible final values for r2 are 6 and 0, depending on whether or not
+P1's store to x propagates to P0 before P0's load from x executes.
+But since P0's load from x is a plain access, the compiler may decide
+to carry out the load twice (for the comparison against NULL, then again
+for the READ_ONCE()) and eliminate the temporary variable r1.  The
+object code generated for P0 could therefore end up looking rather
+like this:
+
+	P0()
+	{
+		int r2 = 0;
+
+		if (x != NULL)
+			r2 = READ_ONCE(*x);
+	}
+
+And now it is obvious that this code runs the risk of dereferencing a
+NULL pointer, because P1's store to x might propagate to P0 after the
+test against NULL has been made but before the READ_ONCE() executes.
+If the original code had said "r1 = READ_ONCE(x)" instead of "r1 = x",
+the compiler would not have performed this optimization and there
+would be no possibility of a NULL-pointer dereference.
+
+Given the possibility of transformations like this one, the LKMM
+doesn't try to predict all possible outcomes of code containing plain
+accesses.  It is instead content to determine whether the code
+violates the compiler's assumptions, which would render the ultimate
+outcome undefined.
+
+In technical terms, the compiler is allowed to assume that when the
+program executes, there will not be any data races.  A "data race"
+occurs when there are two memory accesses such that:
+
+1.	they access the same location,
+
+2.	at least one of them is a store,
+
+3.	at least one of them is plain,
+
+4.	they occur on different CPUs (or in different threads on the
+	same CPU), and
+
+5.	they execute concurrently.
+
+In the literature, two accesses are said to "conflict" if they satisfy
+1 and 2 above.  We'll go a little farther and say that two accesses
+are "race candidates" if they satisfy 1 - 4.  Thus, whether or not two
+race candidates actually do race in a given execution depends on
+whether they are concurrent.
+
+The LKMM tries to determine whether a program contains race candidates
+which may execute concurrently; if it does then the LKMM says there is
+a potential data race and makes no predictions about the program's
+outcome.
+
+Determining whether two accesses are race candidates is easy; you can
+see that all the concepts involved in the definition above are already
+part of the memory model.  The hard part is telling whether they may
+execute concurrently.  The LKMM takes a conservative attitude,
+assuming that accesses may be concurrent unless it can prove they
+are not.
+
+If two memory accesses aren't concurrent then one must execute before
+the other.  Therefore the LKMM decides two accesses aren't concurrent
+if they can be connected by a sequence of hb, pb, and rb links
+(together referred to as xb, for "executes before").  However, there
+are two complicating factors.
+
+If X is a load and X executes before a store Y, then indeed there is
+no danger of X and Y being concurrent.  After all, Y can't have any
+effect on the value obtained by X until the memory subsystem has
+propagated Y from its own CPU to X's CPU, which won't happen until
+some time after Y executes and thus after X executes.  But if X is a
+store, then even if X executes before Y it is still possible that X
+will propagate to Y's CPU just as Y is executing.  In such a case X
+could very well interfere somehow with Y, and we would have to
+consider X and Y to be concurrent.
+
+Therefore when X is a store, for X and Y to be non-concurrent the LKMM
+requires not only that X must execute before Y but also that X must
+propagate to Y's CPU before Y executes.  (Or vice versa, of course, if
+Y executes before X -- then Y must propagate to X's CPU before X
+executes if Y is a store.)  This is expressed by the visibility
+relation (vis), where X ->vis Y is defined to hold if there is an
+intermediate event Z such that:
+
+	X is connected to Z by a possibly empty sequence of
+	cumul-fence links followed by an optional rfe link (if none of
+	these links are present, X and Z are the same event),
+
+and either:
+
+	Z is connected to Y by a strong-fence link followed by a
+	possibly empty sequence of xb links,
+
+or:
+
+	Z is on the same CPU as Y and is connected to Y by a possibly
+	empty sequence of xb links (again, if the sequence is empty it
+	means Z and Y are the same event).
+
+The motivations behind this definition are straightforward:
+
+	cumul-fence memory barriers force stores that are po-before
+	the barrier to propagate to other CPUs before stores that are
+	po-after the barrier.
+
+	An rfe link from an event W to an event R says that R reads
+	from W, which certainly means that W must have propagated to
+	R's CPU before R executed.
+
+	strong-fence memory barriers force stores that are po-before
+	the barrier, or that propagate to the barrier's CPU before the
+	barrier executes, to propagate to all CPUs before any events
+	po-after the barrier can execute.
+
+To see how this works out in practice, consider our old friend, the MP
+pattern (with fences and statement labels, but without the conditional
+test):
+
+	int buf = 0, flag = 0;
+
+	P0()
+	{
+		X: WRITE_ONCE(buf, 1);
+		   smp_wmb();
+		W: WRITE_ONCE(flag, 1);
+	}
+
+	P1()
+	{
+		int r1;
+		int r2 = 0;
+
+		Z: r1 = READ_ONCE(flag);
+		   smp_rmb();
+		Y: r2 = READ_ONCE(buf);
+	}
+
+The smp_wmb() memory barrier gives a cumul-fence link from X to W, and
+assuming r1 = 1 at the end, there is an rfe link from W to Z.  This
+means that the store to buf must propagate from P0 to P1 before Z
+executes.  Next, Z and Y are on the same CPU and the smp_rmb() fence
+provides an xb link from Z to Y (i.e., it forces Z to execute before
+Y).  Therefore we have X ->vis Y: X must propagate to Y's CPU before Y
+executes.
+
+The second complicating factor mentioned above arises from the fact
+that when we are considering data races, some of the memory accesses
+are plain.  Now, although we have not said so explicitly, up to this
+point most of the relations defined by the LKMM (ppo, hb, prop,
+cumul-fence, pb, and so on -- including vis) apply only to marked
+accesses.
+
+There are good reasons for this restriction.  The compiler is not
+allowed to apply fancy transformations to marked accesses, and
+consequently each such access in the source code corresponds more or
+less directly to a single machine instruction in the object code.  But
+plain accesses are a different story; the compiler may combine them,
+split them up, duplicate them, eliminate them, invent new ones, and
+who knows what else.  Seeing a plain access in the source code tells
+you almost nothing about what machine instructions will end up in the
+object code.
+
+Fortunately, the compiler isn't completely free; it is subject to some
+limitations.  For one, it is not allowed to introduce a data race into
+the object code if the source code does not already contain a data
+race (if it could, memory models would be useless and no multithreaded
+code would be safe!).  For another, it cannot move a plain access past
+a compiler barrier.
+
+A compiler barrier is a kind of fence, but as the name implies, it
+only affects the compiler; it does not necessarily have any effect on
+how instructions are executed by the CPU.  In Linux kernel source
+code, the barrier() function is a compiler barrier.  It doesn't give
+rise directly to any machine instructions in the object code; rather,
+it affects how the compiler generates the rest of the object code.
+Given source code like this:
+
+	... some memory accesses ...
+	barrier();
+	... some other memory accesses ...
+
+the barrier() function ensures that the machine instructions
+corresponding to the first group of accesses will all end po-before
+any machine instructions corresponding to the second group of accesses
+-- even if some of the accesses are plain.  (Of course, the CPU may
+then execute some of those accesses out of program order, but we
+already know how to deal with such issues.)  Without the barrier()
+there would be no such guarantee; the two groups of accesses could be
+intermingled or even reversed in the object code.
+
+The LKMM doesn't say much about the barrier() function, but it does
+require that all fences are also compiler barriers.  In addition, it
+requires that the ordering properties of memory barriers such as
+smp_rmb() or smp_store_release() apply to plain accesses as well as to
+marked accesses.
+
+This is the key to analyzing data races.  Consider the MP pattern
+again, now using plain accesses for buf:
+
+	int buf = 0, flag = 0;
+
+	P0()
+	{
+		U: buf = 1;
+		   smp_wmb();
+		X: WRITE_ONCE(flag, 1);
+	}
+
+	P1()
+	{
+		int r1;
+		int r2 = 0;
+
+		Y: r1 = READ_ONCE(flag);
+		   if (r1) {
+			   smp_rmb();
+			V: r2 = buf;
+		   }
+	}
+
+This program does not contain a data race.  Although the U and V
+accesses are race candidates, the LKMM can prove they are not
+concurrent as follows:
+
+	The smp_wmb() fence in P0 is both a compiler barrier and a
+	cumul-fence.  It guarantees that no matter what hash of
+	machine instructions the compiler generates for the plain
+	access U, all those instructions will be po-before the fence.
+	Consequently U's store to buf, no matter how it is carried out
+	at the machine level, must propagate to P1 before X's store to
+	flag does.
+
+	X and Y are both marked accesses.  Hence an rfe link from X to
+	Y is a valid indicator that X propagated to P1 before Y
+	executed, i.e., X ->vis Y.  (And if there is no rfe link then
+	r1 will be 0, so V will not be executed and ipso facto won't
+	race with U.)
+
+	The smp_rmb() fence in P1 is a compiler barrier as well as a
+	fence.  It guarantees that all the machine-level instructions
+	corresponding to the access V will be po-after the fence, and
+	therefore any loads among those instructions will execute
+	after the fence does and hence after Y does.
+
+Thus U's store to buf is forced to propagate to P1 before V's load
+executes (assuming V does execute), ruling out the possibility of a
+data race between them.
+
+This analysis illustrates how the LKMM deals with plain accesses in
+general.  Suppose R is a plain load and we want to show that R
+executes before some marked access E.  We can do this by finding a
+marked access X such that R and X are ordered by a suitable fence and
+X ->xb* E.  If E was also a plain access, we would also look for a
+marked access Y such that X ->xb* Y, and Y and E are ordered by a
+fence.  We describe this arrangement by saying that R is
+"post-bounded" by X and E is "pre-bounded" by Y.
+
+In fact, we go one step further: Since R is a read, we say that R is
+"r-post-bounded" by X.  Similarly, E would be "r-pre-bounded" or
+"w-pre-bounded" by Y, depending on whether E was a store or a load.
+This distinction is needed because some fences affect only loads
+(i.e., smp_rmb()) and some affect only stores (smp_wmb()); otherwise
+the two types of bounds are the same.  And as a degenerate case, we
+say that a marked access pre-bounds and post-bounds itself (e.g., if R
+above were a marked load then X could simply be taken to be R itself.)
+
+The need to distinguish between r- and w-bounding raises yet another
+issue.  When the source code contains a plain store, the compiler is
+allowed to put plain loads of the same location into the object code.
+For example, given the source code:
+
+	x = 1;
+
+the compiler is theoretically allowed to generate object code that
+looks like:
+
+	if (x != 1)
+		x = 1;
+
+thereby adding a load (and possibly replacing the store entirely).
+For this reason, whenever the LKMM requires a plain store to be
+w-pre-bounded or w-post-bounded by a marked access, it also requires
+the store to be r-pre-bounded or r-post-bounded, so as to handle cases
+where the compiler adds a load.
+
+(This may be overly cautious.  We don't know of any examples where a
+compiler has augmented a store with a load in this fashion, and the
+Linux kernel developers would probably fight pretty hard to change a
+compiler if it ever did this.  Still, better safe than sorry.)
+
+Incidentally, the other tranformation -- augmenting a plain load by
+adding in a store to the same location -- is not allowed.  This is
+because the compiler cannot know whether any other CPUs might perform
+a concurrent load from that location.  Two concurrent loads don't
+constitute a race (they can't interfere with each other), but a store
+does race with a concurrent load.  Thus adding a store might create a
+data race where one was not already present in the source code,
+something the compiler is forbidden to do.  Augmenting a store with a
+load, on the other hand, is acceptable because doing so won't create a
+data race unless one already existed.
+
+The LKMM includes a second way to pre-bound plain accesses, in
+addition to fences: an address dependency from a marked load.  That
+is, in the sequence:
+
+	p = READ_ONCE(ptr);
+	r = *p;
+
+the LKMM says that the marked load of ptr pre-bounds the plain load of
+*p; the marked load must execute before any of the machine
+instructions corresponding to the plain load.  This is a reasonable
+stipulation, since after all, the CPU can't perform the load of *p
+until it knows what value p will hold.  Furthermore, without some
+assumption like this one, some usages typical of RCU would count as
+data races.  For example:
+
+	int a = 1, b;
+	int *ptr = &a;
+
+	P0()
+	{
+		b = 2;
+		rcu_assign_pointer(ptr, &b);
+	}
+
+	P1()
+	{
+		int *p;
+		int r;
+
+		rcu_read_lock();
+		p = rcu_dereference(ptr);
+		r = *p;
+		rcu_read_unlock();
+	}
+
+(In this example the rcu_read_lock() and rcu_read_unlock() calls don't
+really do anything, because there aren't any grace periods.  They are
+included merely for the sake of good form; typically P0 would call
+synchronize_rcu() somewhere after the rcu_assign_pointer().)
+
+rcu_assign_pointer() performs a store-release, so the plain store to b
+is definitely w-post-bounded before the store to ptr, and the two
+stores will propagate to P1 in that order.  However, rcu_dereference()
+is only equivalent to READ_ONCE().  While it is a marked access, it is
+not a fence or compiler barrier.  Hence the only guarantee we have
+that the load of ptr in P1 is r-pre-bounded before the load of *p
+(thus avoiding a race) is the assumption about address dependencies.
+
+This is a situation where the compiler can undermine the memory model,
+and a certain amount of care is required when programming constructs
+like this one.  In particular, comparisons between the pointer and
+other known addresses can cause trouble.  If you have something like:
+
+	p = rcu_dereference(ptr);
+	if (p == &x)
+		r = *p;
+
+then the compiler just might generate object code resembling:
+
+	p = rcu_dereference(ptr);
+	if (p == &x)
+		r = x;
+
+or even:
+
+	rtemp = x;
+	p = rcu_dereference(ptr);
+	if (p == &x)
+		r = rtemp;
+
+which would invalidate the memory model's assumption, since the CPU
+could now perform the load of x before the load of ptr (there might be
+a control dependency but no address dependency at the machine level).
+
+Finally, it turns out there is a situation in which a plain write does
+not need to be w-post-bounded: when it is separated from the other
+race-candidate access by a fence.  At first glance this may seem
+impossible.  After all, to be race candidates the two accesses must
+be on different CPUs, and fences don't link events on different CPUs.
+Well, normal fences don't -- but rcu-fence can!  Here's an example:
+
+	int x, y;
+
+	P0()
+	{
+		WRITE_ONCE(x, 1);
+		synchronize_rcu();
+		y = 3;
+	}
+
+	P1()
+	{
+		rcu_read_lock();
+		if (READ_ONCE(x) == 0)
+			y = 2;
+		rcu_read_unlock();
+	}
+
+Do the plain stores to y race?  Clearly not if P1 reads a non-zero
+value for x, so let's assume the READ_ONCE(x) does obtain 0.  This
+means that the read-side critical section in P1 must finish executing
+before the grace period in P0 does, because RCU's Grace-Period
+Guarantee says that otherwise P0's store to x would have propagated to
+P1 before the critical section started and so would have been visible
+to the READ_ONCE().  (Another way of putting it is that the fre link
+from the READ_ONCE() to the WRITE_ONCE() gives rise to an rcu-link
+between those two events.)
+
+This means there is an rcu-fence link from P1's "y = 2" store to P0's
+"y = 3" store, and consequently the first must propagate from P1 to P0
+before the second can execute.  Therefore the two stores cannot be
+concurrent and there is no race, even though P1's plain store to y
+isn't w-post-bounded by any marked accesses.
+
+Putting all this material together yields the following picture.  For
+race-candidate stores W and W', where W ->co W', the LKMM says the
+stores don't race if W can be linked to W' by a
+
+	w-post-bounded ; vis ; w-pre-bounded
+
+sequence.  If W is plain then they also have to be linked by an
+
+	r-post-bounded ; xb* ; w-pre-bounded
+
+sequence, and if W' is plain then they also have to be linked by a
+
+	w-post-bounded ; vis ; r-pre-bounded
+
+sequence.  For race-candidate load R and store W, the LKMM says the
+two accesses don't race if R can be linked to W by an
+
+	r-post-bounded ; xb* ; w-pre-bounded
+
+sequence or if W can be linked to R by a
+
+	w-post-bounded ; vis ; r-pre-bounded
+
+sequence.  For the cases involving a vis link, the LKMM also accepts
+sequences in which W is linked to W' or R by a
+
+	strong-fence ; xb* ; {w and/or r}-pre-bounded
+
+sequence with no post-bounding, and in every case the LKMM also allows
+the link simply to be a fence with no bounding at all.  If no sequence
+of the appropriate sort exists, the LKMM says that the accesses race.
+
+There is one more part of the LKMM related to plain accesses (although
+not to data races) we should discuss.  Recall that many relations such
+as hb are limited to marked accesses only.  As a result, the
+happens-before, propagates-before, and rcu axioms (which state that
+various relation must not contain a cycle) doesn't apply to plain
+accesses.  Nevertheless, we do want to rule out such cycles, because
+they don't make sense even for plain accesses.
+
+To this end, the LKMM imposes three extra restrictions, together
+called the "plain-coherence" axiom because of their resemblance to the
+rules used by the operational model to ensure cache coherence (that
+is, the rules governing the memory subsystem's choice of a store to
+satisfy a load request and its determination of where a store will
+fall in the coherence order):
+
+	If R and W are race candidates and it is possible to link R to
+	W by one of the xb* sequences listed above, then W ->rfe R is
+	not allowed (i.e., a load cannot read from a store that it
+	executes before, even if one or both is plain).
+
+	If W and R are race candidates and it is possible to link W to
+	R by one of the vis sequences listed above, then R ->fre W is
+	not allowed (i.e., if a store is visible to a load then the
+	load must read from that store or one coherence-after it).
+
+	If W and W' are race candidates and it is possible to link W
+	to W' by one of the vis sequences listed above, then W' ->co W
+	is not allowed (i.e., if one store is visible to a second then
+	the second must come after the first in the coherence order).
+
+This is the extent to which the LKMM deals with plain accesses.
+Perhaps it could say more (for example, plain accesses might
+contribute to the ppo relation), but at the moment it seems that this
+minimal, conservative approach is good enough.
+
+
+ODDS AND ENDS
+-------------
+
+This section covers material that didn't quite fit anywhere in the
+earlier sections.
+
+The descriptions in this document don't always match the formal
+version of the LKMM exactly.  For example, the actual formal
+definition of the prop relation makes the initial coe or fre part
+optional, and it doesn't require the events linked by the relation to
+be on the same CPU.  These differences are very unimportant; indeed,
+instances where the coe/fre part of prop is missing are of no interest
+because all the other parts (fences and rfe) are already included in
+hb anyway, and where the formal model adds prop into hb, it includes
+an explicit requirement that the events being linked are on the same
+CPU.
+
+Another minor difference has to do with events that are both memory
+accesses and fences, such as those corresponding to smp_load_acquire()
+calls.  In the formal model, these events aren't actually both reads
+and fences; rather, they are read events with an annotation marking
+them as acquires.  (Or write events annotated as releases, in the case
+smp_store_release().)  The final effect is the same.
+
+Although we didn't mention it above, the instruction execution
+ordering provided by the smp_rmb() fence doesn't apply to read events
+that are part of a non-value-returning atomic update.  For instance,
+given:
+
+	atomic_inc(&x);
+	smp_rmb();
+	r1 = READ_ONCE(y);
+
+it is not guaranteed that the load from y will execute after the
+update to x.  This is because the ARMv8 architecture allows
+non-value-returning atomic operations effectively to be executed off
+the CPU.  Basically, the CPU tells the memory subsystem to increment
+x, and then the increment is carried out by the memory hardware with
+no further involvement from the CPU.  Since the CPU doesn't ever read
+the value of x, there is nothing for the smp_rmb() fence to act on.
+
+The LKMM defines a few extra synchronization operations in terms of
+things we have already covered.  In particular, rcu_dereference() is
+treated as READ_ONCE() and rcu_assign_pointer() is treated as
+smp_store_release() -- which is basically how the Linux kernel treats
+them.
+
+Although we said that plain accesses are not linked by the ppo
+relation, they do contribute to it indirectly.  Firstly, when there is
+an address dependency from a marked load R to a plain store W,
+followed by smp_wmb() and then a marked store W', the LKMM creates a
+ppo link from R to W'.  The reasoning behind this is perhaps a little
+shaky, but essentially it says there is no way to generate object code
+for this source code in which W' could execute before R.  Just as with
+pre-bounding by address dependencies, it is possible for the compiler
+to undermine this relation if sufficient care is not taken.
+
+Secondly, plain accesses can carry dependencies: If a data dependency
+links a marked load R to a store W, and the store is read by a load R'
+from the same thread, then the data loaded by R' depends on the data
+loaded originally by R. Thus, if R' is linked to any access X by a
+dependency, R is also linked to access X by the same dependency, even
+if W' or R' (or both!) are plain.
+
+There are a few oddball fences which need special treatment:
+smp_mb__before_atomic(), smp_mb__after_atomic(), and
+smp_mb__after_spinlock().  The LKMM uses fence events with special
+annotations for them; they act as strong fences just like smp_mb()
+except for the sets of events that they order.  Instead of ordering
+all po-earlier events against all po-later events, as smp_mb() does,
+they behave as follows:
+
+	smp_mb__before_atomic() orders all po-earlier events against
+	po-later atomic updates and the events following them;
+
+	smp_mb__after_atomic() orders po-earlier atomic updates and
+	the events preceding them against all po-later events;
+
+	smp_mb__after_spinlock() orders po-earlier lock acquisition
+	events and the events preceding them against all po-later
+	events.
+
+Interestingly, RCU and locking each introduce the possibility of
+deadlock.  When faced with code sequences such as:
+
+	spin_lock(&s);
+	spin_lock(&s);
+	spin_unlock(&s);
+	spin_unlock(&s);
+
+or:
+
+	rcu_read_lock();
+	synchronize_rcu();
+	rcu_read_unlock();
+
+what does the LKMM have to say?  Answer: It says there are no allowed
+executions at all, which makes sense.  But this can also lead to
+misleading results, because if a piece of code has multiple possible
+executions, some of which deadlock, the model will report only on the
+non-deadlocking executions.  For example:
+
+	int x, y;
+
+	P0()
+	{
+		int r0;
+
+		WRITE_ONCE(x, 1);
+		r0 = READ_ONCE(y);
+	}
+
+	P1()
+	{
+		rcu_read_lock();
+		if (READ_ONCE(x) > 0) {
+			WRITE_ONCE(y, 36);
+			synchronize_rcu();
+		}
+		rcu_read_unlock();
+	}
+
+Is it possible to end up with r0 = 36 at the end?  The LKMM will tell
+you it is not, but the model won't mention that this is because P1
+will self-deadlock in the executions where it stores 36 in y.
diff --git a/tools/memory-model/Documentation/glossary.txt b/tools/memory-model/Documentation/glossary.txt
new file mode 100644
index 0000000000..6f3d16dbf4
--- /dev/null
+++ b/tools/memory-model/Documentation/glossary.txt
@@ -0,0 +1,178 @@
+This document contains brief definitions of LKMM-related terms.  Like most
+glossaries, it is not intended to be read front to back (except perhaps
+as a way of confirming a diagnosis of OCD), but rather to be searched
+for specific terms.
+
+
+Address Dependency:  When the address of a later memory access is computed
+	based on the value returned by an earlier load, an "address
+	dependency" extends from that load extending to the later access.
+	Address dependencies are quite common in RCU read-side critical
+	sections:
+
+	 1 rcu_read_lock();
+	 2 p = rcu_dereference(gp);
+	 3 do_something(p->a);
+	 4 rcu_read_unlock();
+
+	 In this case, because the address of "p->a" on line 3 is computed
+	 from the value returned by the rcu_dereference() on line 2, the
+	 address dependency extends from that rcu_dereference() to that
+	 "p->a".  In rare cases, optimizing compilers can destroy address
+	 dependencies.	Please see Documentation/RCU/rcu_dereference.rst
+	 for more information.
+
+	 See also "Control Dependency" and "Data Dependency".
+
+Acquire:  With respect to a lock, acquiring that lock, for example,
+	using spin_lock().  With respect to a non-lock shared variable,
+	a special operation that includes a load and which orders that
+	load before later memory references running on that same CPU.
+	An example special acquire operation is smp_load_acquire(),
+	but atomic_read_acquire() and atomic_xchg_acquire() also include
+	acquire loads.
+
+	When an acquire load returns the value stored by a release store
+	to that same variable, (in other words, the acquire load "reads
+	from" the release store), then all operations preceding that
+	store "happen before" any operations following that load acquire.
+
+	See also "Happens-Before", "Reads-From", "Relaxed", and "Release".
+
+Coherence (co):  When one CPU's store to a given variable overwrites
+	either the value from another CPU's store or some later value,
+	there is said to be a coherence link from the second CPU to
+	the first.
+
+	It is also possible to have a coherence link within a CPU, which
+	is a "coherence internal" (coi) link.  The term "coherence
+	external" (coe) link is used when it is necessary to exclude
+	the coi case.
+
+	See also "From-reads" and "Reads-from".
+
+Control Dependency:  When a later store's execution depends on a test
+	of a value computed from a value returned by an earlier load,
+	a "control dependency" extends from that load to that store.
+	For example:
+
+	 1 if (READ_ONCE(x))
+	 2   WRITE_ONCE(y, 1);
+
+	 Here, the control dependency extends from the READ_ONCE() on
+	 line 1 to the WRITE_ONCE() on line 2.	Control dependencies are
+	 fragile, and can be easily destroyed by optimizing compilers.
+	 Please see control-dependencies.txt for more information.
+
+	 See also "Address Dependency" and "Data Dependency".
+
+Cycle:	Memory-barrier pairing is restricted to a pair of CPUs, as the
+	name suggests.	And in a great many cases, a pair of CPUs is all
+	that is required.  In other cases, the notion of pairing must be
+	extended to additional CPUs, and the result is called a "cycle".
+	In a cycle, each CPU's ordering interacts with that of the next:
+
+	CPU 0                CPU 1                CPU 2
+	WRITE_ONCE(x, 1);    WRITE_ONCE(y, 1);    WRITE_ONCE(z, 1);
+	smp_mb();            smp_mb();            smp_mb();
+	r0 = READ_ONCE(y);   r1 = READ_ONCE(z);   r2 = READ_ONCE(x);
+
+	CPU 0's smp_mb() interacts with that of CPU 1, which interacts
+	with that of CPU 2, which in turn interacts with that of CPU 0
+	to complete the cycle.	Because of the smp_mb() calls between
+	each pair of memory accesses, the outcome where r0, r1, and r2
+	are all equal to zero is forbidden by LKMM.
+
+	See also "Pairing".
+
+Data Dependency:  When the data written by a later store is computed based
+	on the value returned by an earlier load, a "data dependency"
+	extends from that load to that later store.  For example:
+
+	 1 r1 = READ_ONCE(x);
+	 2 WRITE_ONCE(y, r1 + 1);
+
+	In this case, the data dependency extends from the READ_ONCE()
+	on line 1 to the WRITE_ONCE() on line 2.  Data dependencies are
+	fragile and can be easily destroyed by optimizing compilers.
+	Because optimizing compilers put a great deal of effort into
+	working out what values integer variables might have, this is
+	especially true in cases where the dependency is carried through
+	an integer.
+
+	See also "Address Dependency" and "Control Dependency".
+
+From-Reads (fr):  When one CPU's store to a given variable happened
+	too late to affect the value returned by another CPU's
+	load from that same variable, there is said to be a from-reads
+	link from the load to the store.
+
+	It is also possible to have a from-reads link within a CPU, which
+	is a "from-reads internal" (fri) link.  The term "from-reads
+	external" (fre) link is used when it is necessary to exclude
+	the fri case.
+
+	See also "Coherence" and "Reads-from".
+
+Fully Ordered:  An operation such as smp_mb() that orders all of
+	its CPU's prior accesses with all of that CPU's subsequent
+	accesses, or a marked access such as atomic_add_return()
+	that orders all of its CPU's prior accesses, itself, and
+	all of its CPU's subsequent accesses.
+
+Happens-Before (hb): A relation between two accesses in which LKMM
+	guarantees the first access precedes the second.  For more
+	detail, please see the "THE HAPPENS-BEFORE RELATION: hb"
+	section of explanation.txt.
+
+Marked Access:  An access to a variable that uses an special function or
+	macro such as "r1 = READ_ONCE(x)" or "smp_store_release(&a, 1)".
+
+	See also "Unmarked Access".
+
+Pairing: "Memory-barrier pairing" reflects the fact that synchronizing
+	data between two CPUs requires that both CPUs their accesses.
+	Memory barriers thus tend to come in pairs, one executed by
+	one of the CPUs and the other by the other CPU.  Of course,
+	pairing also occurs with other types of operations, so that a
+	smp_store_release() pairs with an smp_load_acquire() that reads
+	the value stored.
+
+	See also "Cycle".
+
+Reads-From (rf):  When one CPU's load returns the value stored by some other
+	CPU, there is said to be a reads-from link from the second
+	CPU's store to the first CPU's load.  Reads-from links have the
+	nice property that time must advance from the store to the load,
+	which means that algorithms using reads-from links can use lighter
+	weight ordering and synchronization compared to algorithms using
+	coherence and from-reads links.
+
+	It is also possible to have a reads-from link within a CPU, which
+	is a "reads-from internal" (rfi) link.	The term "reads-from
+	external" (rfe) link is used when it is necessary to exclude
+	the rfi case.
+
+	See also Coherence" and "From-reads".
+
+Relaxed:  A marked access that does not imply ordering, for example, a
+	READ_ONCE(), WRITE_ONCE(), a non-value-returning read-modify-write
+	operation, or a value-returning read-modify-write operation whose
+	name ends in "_relaxed".
+
+	See also "Acquire" and "Release".
+
+Release:  With respect to a lock, releasing that lock, for example,
+	using spin_unlock().  With respect to a non-lock shared variable,
+	a special operation that includes a store and which orders that
+	store after earlier memory references that ran on that same CPU.
+	An example special release store is smp_store_release(), but
+	atomic_set_release() and atomic_cmpxchg_release() also include
+	release stores.
+
+	See also "Acquire" and "Relaxed".
+
+Unmarked Access:  An access to a variable that uses normal C-language
+	syntax, for example, "a = b[2]";
+
+	See also "Marked Access".
diff --git a/tools/memory-model/Documentation/litmus-tests.txt b/tools/memory-model/Documentation/litmus-tests.txt
new file mode 100644
index 0000000000..acac527328
--- /dev/null
+++ b/tools/memory-model/Documentation/litmus-tests.txt
@@ -0,0 +1,1083 @@
+Linux-Kernel Memory Model Litmus Tests
+======================================
+
+This file describes the LKMM litmus-test format by example, describes
+some tricks and traps, and finally outlines LKMM's limitations.  Earlier
+versions of this material appeared in a number of LWN articles, including:
+
+https://lwn.net/Articles/720550/
+	A formal kernel memory-ordering model (part 2)
+https://lwn.net/Articles/608550/
+	Axiomatic validation of memory barriers and atomic instructions
+https://lwn.net/Articles/470681/
+	Validating Memory Barriers and Atomic Instructions
+
+This document presents information in decreasing order of applicability,
+so that, where possible, the information that has proven more commonly
+useful is shown near the beginning.
+
+For information on installing LKMM, including the underlying "herd7"
+tool, please see tools/memory-model/README.
+
+
+Copy-Pasta
+==========
+
+As with other software, it is often better (if less macho) to adapt an
+existing litmus test than it is to create one from scratch.  A number
+of litmus tests may be found in the kernel source tree:
+
+	tools/memory-model/litmus-tests/
+	Documentation/litmus-tests/
+
+Several thousand more example litmus tests are available on github
+and kernel.org:
+
+	https://github.com/paulmckrcu/litmus
+	https://git.kernel.org/pub/scm/linux/kernel/git/paulmck/perfbook.git/tree/CodeSamples/formal/herd
+	https://git.kernel.org/pub/scm/linux/kernel/git/paulmck/perfbook.git/tree/CodeSamples/formal/litmus
+
+The -l and -L arguments to "git grep" can be quite helpful in identifying
+existing litmus tests that are similar to the one you need.  But even if
+you start with an existing litmus test, it is still helpful to have a
+good understanding of the litmus-test format.
+
+
+Examples and Format
+===================
+
+This section describes the overall format of litmus tests, starting
+with a small example of the message-passing pattern and moving on to
+more complex examples that illustrate explicit initialization and LKMM's
+minimalistic set of flow-control statements.
+
+
+Message-Passing Example
+-----------------------
+
+This section gives an overview of the format of a litmus test using an
+example based on the common message-passing use case.  This use case
+appears often in the Linux kernel.  For example, a flag (modeled by "y"
+below) indicates that a buffer (modeled by "x" below) is now completely
+filled in and ready for use.  It would be very bad if the consumer saw the
+flag set, but, due to memory misordering, saw old values in the buffer.
+
+This example asks whether smp_store_release() and smp_load_acquire()
+suffices to avoid this bad outcome:
+
+ 1 C MP+pooncerelease+poacquireonce
+ 2
+ 3 {}
+ 4
+ 5 P0(int *x, int *y)
+ 6 {
+ 7   WRITE_ONCE(*x, 1);
+ 8   smp_store_release(y, 1);
+ 9 }
+10
+11 P1(int *x, int *y)
+12 {
+13   int r0;
+14   int r1;
+15
+16   r0 = smp_load_acquire(y);
+17   r1 = READ_ONCE(*x);
+18 }
+19
+20 exists (1:r0=1 /\ 1:r1=0)
+
+Line 1 starts with "C", which identifies this file as being in the
+LKMM C-language format (which, as we will see, is a small fragment
+of the full C language).  The remainder of line 1 is the name of
+the test, which by convention is the filename with the ".litmus"
+suffix stripped.  In this case, the actual test may be found in
+tools/memory-model/litmus-tests/MP+pooncerelease+poacquireonce.litmus
+in the Linux-kernel source tree.
+
+Mechanically generated litmus tests will often have an optional
+double-quoted comment string on the second line.  Such strings are ignored
+when running the test.  Yes, you can add your own comments to litmus
+tests, but this is a bit involved due to the use of multiple parsers.
+For now, you can use C-language comments in the C code, and these comments
+may be in either the "/* */" or the "//" style.  A later section will
+cover the full litmus-test commenting story.
+
+Line 3 is the initialization section.  Because the default initialization
+to zero suffices for this test, the "{}" syntax is used, which mean the
+initialization section is empty.  Litmus tests requiring non-default
+initialization must have non-empty initialization sections, as in the
+example that will be presented later in this document.
+
+Lines 5-9 show the first process and lines 11-18 the second process.  Each
+process corresponds to a Linux-kernel task (or kthread, workqueue, thread,
+and so on; LKMM discussions often use these terms interchangeably).
+The name of the first process is "P0" and that of the second "P1".
+You can name your processes anything you like as long as the names consist
+of a single "P" followed by a number, and as long as the numbers are
+consecutive starting with zero.  This can actually be quite helpful,
+for example, a .litmus file matching "^P1(" but not matching "^P2("
+must contain a two-process litmus test.
+
+The argument list for each function are pointers to the global variables
+used by that function.  Unlike normal C-language function parameters, the
+names are significant.  The fact that both P0() and P1() have a formal
+parameter named "x" means that these two processes are working with the
+same global variable, also named "x".  So the "int *x, int *y" on P0()
+and P1() mean that both processes are working with two shared global
+variables, "x" and "y".  Global variables are always passed to processes
+by reference, hence "P0(int *x, int *y)", but *never* "P0(int x, int y)".
+
+P0() has no local variables, but P1() has two of them named "r0" and "r1".
+These names may be freely chosen, but for historical reasons stemming from
+other litmus-test formats, it is conventional to use names consisting of
+"r" followed by a number as shown here.  A common bug in litmus tests
+is forgetting to add a global variable to a process's parameter list.
+This will sometimes result in an error message, but can also cause the
+intended global to instead be silently treated as an undeclared local
+variable.
+
+Each process's code is similar to Linux-kernel C, as can be seen on lines
+7-8 and 13-17.  This code may use many of the Linux kernel's atomic
+operations, some of its exclusive-lock functions, and some of its RCU
+and SRCU functions.  An approximate list of the currently supported
+functions may be found in the linux-kernel.def file.
+
+The P0() process does "WRITE_ONCE(*x, 1)" on line 7.  Because "x" is a
+pointer in P0()'s parameter list, this does an unordered store to global
+variable "x".  Line 8 does "smp_store_release(y, 1)", and because "y"
+is also in P0()'s parameter list, this does a release store to global
+variable "y".
+
+The P1() process declares two local variables on lines 13 and 14.
+Line 16 does "r0 = smp_load_acquire(y)" which does an acquire load
+from global variable "y" into local variable "r0".  Line 17 does a
+"r1 = READ_ONCE(*x)", which does an unordered load from "*x" into local
+variable "r1".  Both "x" and "y" are in P1()'s parameter list, so both
+reference the same global variables that are used by P0().
+
+Line 20 is the "exists" assertion expression to evaluate the final state.
+This final state is evaluated after the dust has settled: both processes
+have completed and all of their memory references and memory barriers
+have propagated to all parts of the system.  The references to the local
+variables "r0" and "r1" in line 24 must be prefixed with "1:" to specify
+which process they are local to.
+
+Note that the assertion expression is written in the litmus-test
+language rather than in C.  For example, single "=" is an equality
+operator rather than an assignment.  The "/\" character combination means
+"and".  Similarly, "\/" stands for "or".  Both of these are ASCII-art
+representations of the corresponding mathematical symbols.  Finally,
+"~" stands for "logical not", which is "!" in C, and not to be confused
+with the C-language "~" operator which instead stands for "bitwise not".
+Parentheses may be used to override precedence.
+
+The "exists" assertion on line 20 is satisfied if the consumer sees the
+flag ("y") set but the buffer ("x") as not yet filled in, that is, if P1()
+loaded a value from "x" that was equal to 1 but loaded a value from "y"
+that was still equal to zero.
+
+This example can be checked by running the following command, which
+absolutely must be run from the tools/memory-model directory and from
+this directory only:
+
+herd7 -conf linux-kernel.cfg litmus-tests/MP+pooncerelease+poacquireonce.litmus
+
+The output is the result of something similar to a full state-space
+search, and is as follows:
+
+ 1 Test MP+pooncerelease+poacquireonce Allowed
+ 2 States 3
+ 3 1:r0=0; 1:r1=0;
+ 4 1:r0=0; 1:r1=1;
+ 5 1:r0=1; 1:r1=1;
+ 6 No
+ 7 Witnesses
+ 8 Positive: 0 Negative: 3
+ 9 Condition exists (1:r0=1 /\ 1:r1=0)
+10 Observation MP+pooncerelease+poacquireonce Never 0 3
+11 Time MP+pooncerelease+poacquireonce 0.00
+12 Hash=579aaa14d8c35a39429b02e698241d09
+
+The most pertinent line is line 10, which contains "Never 0 3", which
+indicates that the bad result flagged by the "exists" clause never
+happens.  This line might instead say "Sometimes" to indicate that the
+bad result happened in some but not all executions, or it might say
+"Always" to indicate that the bad result happened in all executions.
+(The herd7 tool doesn't judge, so it is only an LKMM convention that the
+"exists" clause indicates a bad result.  To see this, invert the "exists"
+clause's condition and run the test.)  The numbers ("0 3") at the end
+of this line indicate the number of end states satisfying the "exists"
+clause (0) and the number not not satisfying that clause (3).
+
+Another important part of this output is shown in lines 2-5, repeated here:
+
+ 2 States 3
+ 3 1:r0=0; 1:r1=0;
+ 4 1:r0=0; 1:r1=1;
+ 5 1:r0=1; 1:r1=1;
+
+Line 2 gives the total number of end states, and each of lines 3-5 list
+one of these states, with the first ("1:r0=0; 1:r1=0;") indicating that
+both of P1()'s loads returned the value "0".  As expected, given the
+"Never" on line 10, the state flagged by the "exists" clause is not
+listed.  This full list of states can be helpful when debugging a new
+litmus test.
+
+The rest of the output is not normally needed, either due to irrelevance
+or due to being redundant with the lines discussed above.  However, the
+following paragraph lists them for the benefit of readers possessed of
+an insatiable curiosity.  Other readers should feel free to skip ahead.
+
+Line 1 echos the test name, along with the "Test" and "Allowed".  Line 6's
+"No" says that the "exists" clause was not satisfied by any execution,
+and as such it has the same meaning as line 10's "Never".  Line 7 is a
+lead-in to line 8's "Positive: 0 Negative: 3", which lists the number
+of end states satisfying and not satisfying the "exists" clause, just
+like the two numbers at the end of line 10.  Line 9 repeats the "exists"
+clause so that you don't have to look it up in the litmus-test file.
+The number at the end of line 11 (which begins with "Time") gives the
+time in seconds required to analyze the litmus test.  Small tests such
+as this one complete in a few milliseconds, so "0.00" is quite common.
+Line 12 gives a hash of the contents for the litmus-test file, and is used
+by tooling that manages litmus tests and their output.  This tooling is
+used by people modifying LKMM itself, and among other things lets such
+people know which of the several thousand relevant litmus tests were
+affected by a given change to LKMM.
+
+
+Initialization
+--------------
+
+The previous example relied on the default zero initialization for
+"x" and "y", but a similar litmus test could instead initialize them
+to some other value:
+
+ 1 C MP+pooncerelease+poacquireonce
+ 2
+ 3 {
+ 4   x=42;
+ 5   y=42;
+ 6 }
+ 7
+ 8 P0(int *x, int *y)
+ 9 {
+10   WRITE_ONCE(*x, 1);
+11   smp_store_release(y, 1);
+12 }
+13
+14 P1(int *x, int *y)
+15 {
+16   int r0;
+17   int r1;
+18
+19   r0 = smp_load_acquire(y);
+20   r1 = READ_ONCE(*x);
+21 }
+22
+23 exists (1:r0=1 /\ 1:r1=42)
+
+Lines 3-6 now initialize both "x" and "y" to the value 42.  This also
+means that the "exists" clause on line 23 must change "1:r1=0" to
+"1:r1=42".
+
+Running the test gives the same overall result as before, but with the
+value 42 appearing in place of the value zero:
+
+ 1 Test MP+pooncerelease+poacquireonce Allowed
+ 2 States 3
+ 3 1:r0=1; 1:r1=1;
+ 4 1:r0=42; 1:r1=1;
+ 5 1:r0=42; 1:r1=42;
+ 6 No
+ 7 Witnesses
+ 8 Positive: 0 Negative: 3
+ 9 Condition exists (1:r0=1 /\ 1:r1=42)
+10 Observation MP+pooncerelease+poacquireonce Never 0 3
+11 Time MP+pooncerelease+poacquireonce 0.02
+12 Hash=ab9a9b7940a75a792266be279a980156
+
+It is tempting to avoid the open-coded repetitions of the value "42"
+by defining another global variable "initval=42" and replacing all
+occurrences of "42" with "initval".  This will not, repeat *not*,
+initialize "x" and "y" to 42, but instead to the address of "initval"
+(try it!).  See the section below on linked lists to learn more about
+why this approach to initialization can be useful.
+
+
+Control Structures
+------------------
+
+LKMM supports the C-language "if" statement, which allows modeling of
+conditional branches.  In LKMM, conditional branches can affect ordering,
+but only if you are *very* careful (compilers are surprisingly able
+to optimize away conditional branches).  The following example shows
+the "load buffering" (LB) use case that is used in the Linux kernel to
+synchronize between ring-buffer producers and consumers.  In the example
+below, P0() is one side checking to see if an operation may proceed and
+P1() is the other side completing its update.
+
+ 1 C LB+fencembonceonce+ctrlonceonce
+ 2
+ 3 {}
+ 4
+ 5 P0(int *x, int *y)
+ 6 {
+ 7   int r0;
+ 8
+ 9   r0 = READ_ONCE(*x);
+10   if (r0)
+11     WRITE_ONCE(*y, 1);
+12 }
+13
+14 P1(int *x, int *y)
+15 {
+16   int r0;
+17
+18   r0 = READ_ONCE(*y);
+19   smp_mb();
+20   WRITE_ONCE(*x, 1);
+21 }
+22
+23 exists (0:r0=1 /\ 1:r0=1)
+
+P1()'s "if" statement on line 10 works as expected, so that line 11 is
+executed only if line 9 loads a non-zero value from "x".  Because P1()'s
+write of "1" to "x" happens only after P1()'s read from "y", one would
+hope that the "exists" clause cannot be satisfied.  LKMM agrees:
+
+ 1 Test LB+fencembonceonce+ctrlonceonce Allowed
+ 2 States 2
+ 3 0:r0=0; 1:r0=0;
+ 4 0:r0=1; 1:r0=0;
+ 5 No
+ 6 Witnesses
+ 7 Positive: 0 Negative: 2
+ 8 Condition exists (0:r0=1 /\ 1:r0=1)
+ 9 Observation LB+fencembonceonce+ctrlonceonce Never 0 2
+10 Time LB+fencembonceonce+ctrlonceonce 0.00
+11 Hash=e5260556f6de495fd39b556d1b831c3b
+
+However, there is no "while" statement due to the fact that full
+state-space search has some difficulty with iteration.  However, there
+are tricks that may be used to handle some special cases, which are
+discussed below.  In addition, loop-unrolling tricks may be applied,
+albeit sparingly.
+
+
+Tricks and Traps
+================
+
+This section covers extracting debug output from herd7, emulating
+spin loops, handling trivial linked lists, adding comments to litmus tests,
+emulating call_rcu(), and finally tricks to improve herd7 performance
+in order to better handle large litmus tests.
+
+
+Debug Output
+------------
+
+By default, the herd7 state output includes all variables mentioned
+in the "exists" clause.  But sometimes debugging efforts are greatly
+aided by the values of other variables.  Consider this litmus test
+(tools/memory-order/litmus-tests/SB+rfionceonce-poonceonces.litmus but
+slightly modified), which probes an obscure corner of hardware memory
+ordering:
+
+ 1 C SB+rfionceonce-poonceonces
+ 2
+ 3 {}
+ 4
+ 5 P0(int *x, int *y)
+ 6 {
+ 7   int r1;
+ 8   int r2;
+ 9
+10   WRITE_ONCE(*x, 1);
+11   r1 = READ_ONCE(*x);
+12   r2 = READ_ONCE(*y);
+13 }
+14
+15 P1(int *x, int *y)
+16 {
+17   int r3;
+18   int r4;
+19
+20   WRITE_ONCE(*y, 1);
+21   r3 = READ_ONCE(*y);
+22   r4 = READ_ONCE(*x);
+23 }
+24
+25 exists (0:r2=0 /\ 1:r4=0)
+
+The herd7 output is as follows:
+
+ 1 Test SB+rfionceonce-poonceonces Allowed
+ 2 States 4
+ 3 0:r2=0; 1:r4=0;
+ 4 0:r2=0; 1:r4=1;
+ 5 0:r2=1; 1:r4=0;
+ 6 0:r2=1; 1:r4=1;
+ 7 Ok
+ 8 Witnesses
+ 9 Positive: 1 Negative: 3
+10 Condition exists (0:r2=0 /\ 1:r4=0)
+11 Observation SB+rfionceonce-poonceonces Sometimes 1 3
+12 Time SB+rfionceonce-poonceonces 0.01
+13 Hash=c7f30fe0faebb7d565405d55b7318ada
+
+(This output indicates that CPUs are permitted to "snoop their own
+store buffers", which all of Linux's CPU families other than s390 will
+happily do.  Such snooping results in disagreement among CPUs on the
+order of stores from different CPUs, which is rarely an issue.)
+
+But the herd7 output shows only the two variables mentioned in the
+"exists" clause.  Someone modifying this test might wish to know the
+values of "x", "y", "0:r1", and "0:r3" as well.  The "locations"
+statement on line 25 shows how to cause herd7 to display additional
+variables:
+
+ 1 C SB+rfionceonce-poonceonces
+ 2
+ 3 {}
+ 4
+ 5 P0(int *x, int *y)
+ 6 {
+ 7   int r1;
+ 8   int r2;
+ 9
+10   WRITE_ONCE(*x, 1);
+11   r1 = READ_ONCE(*x);
+12   r2 = READ_ONCE(*y);
+13 }
+14
+15 P1(int *x, int *y)
+16 {
+17   int r3;
+18   int r4;
+19
+20   WRITE_ONCE(*y, 1);
+21   r3 = READ_ONCE(*y);
+22   r4 = READ_ONCE(*x);
+23 }
+24
+25 locations [0:r1; 1:r3; x; y]
+26 exists (0:r2=0 /\ 1:r4=0)
+
+The herd7 output then displays the values of all the variables:
+
+ 1 Test SB+rfionceonce-poonceonces Allowed
+ 2 States 4
+ 3 0:r1=1; 0:r2=0; 1:r3=1; 1:r4=0; x=1; y=1;
+ 4 0:r1=1; 0:r2=0; 1:r3=1; 1:r4=1; x=1; y=1;
+ 5 0:r1=1; 0:r2=1; 1:r3=1; 1:r4=0; x=1; y=1;
+ 6 0:r1=1; 0:r2=1; 1:r3=1; 1:r4=1; x=1; y=1;
+ 7 Ok
+ 8 Witnesses
+ 9 Positive: 1 Negative: 3
+10 Condition exists (0:r2=0 /\ 1:r4=0)
+11 Observation SB+rfionceonce-poonceonces Sometimes 1 3
+12 Time SB+rfionceonce-poonceonces 0.01
+13 Hash=40de8418c4b395388f6501cafd1ed38d
+
+What if you would like to know the value of a particular global variable
+at some particular point in a given process's execution?  One approach
+is to use a READ_ONCE() to load that global variable into a new local
+variable, then add that local variable to the "locations" clause.
+But be careful:  In some litmus tests, adding a READ_ONCE() will change
+the outcome!  For one example, please see the C-READ_ONCE.litmus and
+C-READ_ONCE-omitted.litmus tests located here:
+
+	https://github.com/paulmckrcu/litmus/blob/master/manual/kernel/
+
+
+Spin Loops
+----------
+
+The analysis carried out by herd7 explores full state space, which is
+at best of exponential time complexity.  Adding processes and increasing
+the amount of code in a give process can greatly increase execution time.
+Potentially infinite loops, such as those used to wait for locks to
+become available, are clearly problematic.
+
+Fortunately, it is possible to avoid state-space explosion by specially
+modeling such loops.  For example, the following litmus tests emulates
+locking using xchg_acquire(), but instead of enclosing xchg_acquire()
+in a spin loop, it instead excludes executions that fail to acquire the
+lock using a herd7 "filter" clause.  Note that for exclusive locking, you
+are better off using the spin_lock() and spin_unlock() that LKMM directly
+models, if for no other reason that these are much faster.  However, the
+techniques illustrated in this section can be used for other purposes,
+such as emulating reader-writer locking, which LKMM does not yet model.
+
+ 1 C C-SB+l-o-o-u+l-o-o-u-X
+ 2
+ 3 {
+ 4 }
+ 5
+ 6 P0(int *sl, int *x0, int *x1)
+ 7 {
+ 8   int r2;
+ 9   int r1;
+10
+11   r2 = xchg_acquire(sl, 1);
+12   WRITE_ONCE(*x0, 1);
+13   r1 = READ_ONCE(*x1);
+14   smp_store_release(sl, 0);
+15 }
+16
+17 P1(int *sl, int *x0, int *x1)
+18 {
+19   int r2;
+20   int r1;
+21
+22   r2 = xchg_acquire(sl, 1);
+23   WRITE_ONCE(*x1, 1);
+24   r1 = READ_ONCE(*x0);
+25   smp_store_release(sl, 0);
+26 }
+27
+28 filter (0:r2=0 /\ 1:r2=0)
+29 exists (0:r1=0 /\ 1:r1=0)
+
+This litmus test may be found here:
+
+https://git.kernel.org/pub/scm/linux/kernel/git/paulmck/perfbook.git/tree/CodeSamples/formal/herd/C-SB+l-o-o-u+l-o-o-u-X.litmus
+
+This test uses two global variables, "x1" and "x2", and also emulates a
+single global spinlock named "sl".  This spinlock is held by whichever
+process changes the value of "sl" from "0" to "1", and is released when
+that process sets "sl" back to "0".  P0()'s lock acquisition is emulated
+on line 11 using xchg_acquire(), which unconditionally stores the value
+"1" to "sl" and stores either "0" or "1" to "r2", depending on whether
+the lock acquisition was successful or unsuccessful (due to "sl" already
+having the value "1"), respectively.  P1() operates in a similar manner.
+
+Rather unconventionally, execution appears to proceed to the critical
+section on lines 12 and 13 in either case.  Line 14 then uses an
+smp_store_release() to store zero to "sl", thus emulating lock release.
+
+The case where xchg_acquire() fails to acquire the lock is handled by
+the "filter" clause on line 28, which tells herd7 to keep only those
+executions in which both "0:r2" and "1:r2" are zero, that is to pay
+attention only to those executions in which both locks are actually
+acquired.  Thus, the bogus executions that would execute the critical
+sections are discarded and any effects that they might have had are
+ignored.  Note well that the "filter" clause keeps those executions
+for which its expression is satisfied, that is, for which the expression
+evaluates to true.  In other words, the "filter" clause says what to
+keep, not what to discard.
+
+The result of running this test is as follows:
+
+ 1 Test C-SB+l-o-o-u+l-o-o-u-X Allowed
+ 2 States 2
+ 3 0:r1=0; 1:r1=1;
+ 4 0:r1=1; 1:r1=0;
+ 5 No
+ 6 Witnesses
+ 7 Positive: 0 Negative: 2
+ 8 Condition exists (0:r1=0 /\ 1:r1=0)
+ 9 Observation C-SB+l-o-o-u+l-o-o-u-X Never 0 2
+10 Time C-SB+l-o-o-u+l-o-o-u-X 0.03
+
+The "Never" on line 9 indicates that this use of xchg_acquire() and
+smp_store_release() really does correctly emulate locking.
+
+Why doesn't the litmus test take the simpler approach of using a spin loop
+to handle failed spinlock acquisitions, like the kernel does?  The key
+insight behind this litmus test is that spin loops have no effect on the
+possible "exists"-clause outcomes of program execution in the absence
+of deadlock.  In other words, given a high-quality lock-acquisition
+primitive in a deadlock-free program running on high-quality hardware,
+each lock acquisition will eventually succeed.  Because herd7 already
+explores the full state space, the length of time required to actually
+acquire the lock does not matter.  After all, herd7 already models all
+possible durations of the xchg_acquire() statements.
+
+Why not just add the "filter" clause to the "exists" clause, thus
+avoiding the "filter" clause entirely?  This does work, but is slower.
+The reason that the "filter" clause is faster is that (in the common case)
+herd7 knows to abandon an execution as soon as the "filter" expression
+fails to be satisfied.  In contrast, the "exists" clause is evaluated
+only at the end of time, thus requiring herd7 to waste time on bogus
+executions in which both critical sections proceed concurrently.  In
+addition, some LKMM users like the separation of concerns provided by
+using the both the "filter" and "exists" clauses.
+
+Readers lacking a pathological interest in odd corner cases should feel
+free to skip the remainder of this section.
+
+But what if the litmus test were to temporarily set "0:r2" to a non-zero
+value?  Wouldn't that cause herd7 to abandon the execution prematurely
+due to an early mismatch of the "filter" clause?
+
+Why not just try it?  Line 4 of the following modified litmus test
+introduces a new global variable "x2" that is initialized to "1".  Line 23
+of P1() reads that variable into "1:r2" to force an early mismatch with
+the "filter" clause.  Line 24 does a known-true "if" condition to avoid
+and static analysis that herd7 might do.  Finally the "exists" clause
+on line 32 is updated to a condition that is alway satisfied at the end
+of the test.
+
+ 1 C C-SB+l-o-o-u+l-o-o-u-X
+ 2
+ 3 {
+ 4   x2=1;
+ 5 }
+ 6
+ 7 P0(int *sl, int *x0, int *x1)
+ 8 {
+ 9   int r2;
+10   int r1;
+11
+12   r2 = xchg_acquire(sl, 1);
+13   WRITE_ONCE(*x0, 1);
+14   r1 = READ_ONCE(*x1);
+15   smp_store_release(sl, 0);
+16 }
+17
+18 P1(int *sl, int *x0, int *x1, int *x2)
+19 {
+20   int r2;
+21   int r1;
+22
+23   r2 = READ_ONCE(*x2);
+24   if (r2)
+25     r2 = xchg_acquire(sl, 1);
+26   WRITE_ONCE(*x1, 1);
+27   r1 = READ_ONCE(*x0);
+28   smp_store_release(sl, 0);
+29 }
+30
+31 filter (0:r2=0 /\ 1:r2=0)
+32 exists (x1=1)
+
+If the "filter" clause were to check each variable at each point in the
+execution, running this litmus test would display no executions because
+all executions would be filtered out at line 23.  However, the output
+is instead as follows:
+
+ 1 Test C-SB+l-o-o-u+l-o-o-u-X Allowed
+ 2 States 1
+ 3 x1=1;
+ 4 Ok
+ 5 Witnesses
+ 6 Positive: 2 Negative: 0
+ 7 Condition exists (x1=1)
+ 8 Observation C-SB+l-o-o-u+l-o-o-u-X Always 2 0
+ 9 Time C-SB+l-o-o-u+l-o-o-u-X 0.04
+10 Hash=080bc508da7f291e122c6de76c0088e3
+
+Line 3 shows that there is one execution that did not get filtered out,
+so the "filter" clause is evaluated only on the last assignment to
+the variables that it checks.  In this case, the "filter" clause is a
+disjunction, so it might be evaluated twice, once at the final (and only)
+assignment to "0:r2" and once at the final assignment to "1:r2".
+
+
+Linked Lists
+------------
+
+LKMM can handle linked lists, but only linked lists in which each node
+contains nothing except a pointer to the next node in the list.  This is
+of course quite restrictive, but there is nevertheless quite a bit that
+can be done within these confines, as can be seen in the litmus test
+at tools/memory-model/litmus-tests/MP+onceassign+derefonce.litmus:
+
+ 1 C MP+onceassign+derefonce
+ 2
+ 3 {
+ 4 y=z;
+ 5 z=0;
+ 6 }
+ 7
+ 8 P0(int *x, int **y)
+ 9 {
+10   WRITE_ONCE(*x, 1);
+11   rcu_assign_pointer(*y, x);
+12 }
+13
+14 P1(int *x, int **y)
+15 {
+16   int *r0;
+17   int r1;
+18
+19   rcu_read_lock();
+20   r0 = rcu_dereference(*y);
+21   r1 = READ_ONCE(*r0);
+22   rcu_read_unlock();
+23 }
+24
+25 exists (1:r0=x /\ 1:r1=0)
+
+Line 4's "y=z" may seem odd, given that "z" has not yet been initialized.
+But "y=z" does not set the value of "y" to that of "z", but instead
+sets the value of "y" to the *address* of "z".  Lines 4 and 5 therefore
+create a simple linked list, with "y" pointing to "z" and "z" having a
+NULL pointer.  A much longer linked list could be created if desired,
+and circular singly linked lists can also be created and manipulated.
+
+The "exists" clause works the same way, with the "1:r0=x" comparing P1()'s
+"r0" not to the value of "x", but again to its address.  This term of the
+"exists" clause therefore tests whether line 20's load from "y" saw the
+value stored by line 11, which is in fact what is required in this case.
+
+P0()'s line 10 initializes "x" to the value 1 then line 11 links to "x"
+from "y", replacing "z".
+
+P1()'s line 20 loads a pointer from "y", and line 21 dereferences that
+pointer.  The RCU read-side critical section spanning lines 19-22 is just
+for show in this example.  Note that the address used for line 21's load
+depends on (in this case, "is exactly the same as") the value loaded by
+line 20.  This is an example of what is called an "address dependency".
+This particular address dependency extends from the load on line 20 to the
+load on line 21.  Address dependencies provide a weak form of ordering.
+
+Running this test results in the following:
+
+ 1 Test MP+onceassign+derefonce Allowed
+ 2 States 2
+ 3 1:r0=x; 1:r1=1;
+ 4 1:r0=z; 1:r1=0;
+ 5 No
+ 6 Witnesses
+ 7 Positive: 0 Negative: 2
+ 8 Condition exists (1:r0=x /\ 1:r1=0)
+ 9 Observation MP+onceassign+derefonce Never 0 2
+10 Time MP+onceassign+derefonce 0.00
+11 Hash=49ef7a741563570102448a256a0c8568
+
+The only possible outcomes feature P1() loading a pointer to "z"
+(which contains zero) on the one hand and P1() loading a pointer to "x"
+(which contains the value one) on the other.  This should be reassuring
+because it says that RCU readers cannot see the old preinitialization
+values when accessing a newly inserted list node.  This undesirable
+scenario is flagged by the "exists" clause, and would occur if P1()
+loaded a pointer to "x", but obtained the pre-initialization value of
+zero after dereferencing that pointer.
+
+
+Comments
+--------
+
+Different portions of a litmus test are processed by different parsers,
+which has the charming effect of requiring different comment syntax in
+different portions of the litmus test.  The C-syntax portions use
+C-language comments (either "/* */" or "//"), while the other portions
+use Ocaml comments "(* *)".
+
+The following litmus test illustrates the comment style corresponding
+to each syntactic unit of the test:
+
+ 1 C MP+onceassign+derefonce (* A *)
+ 2
+ 3 (* B *)
+ 4
+ 5 {
+ 6 y=z; (* C *)
+ 7 z=0;
+ 8 } // D
+ 9
+10 // E
+11
+12 P0(int *x, int **y) // F
+13 {
+14   WRITE_ONCE(*x, 1);  // G
+15   rcu_assign_pointer(*y, x);
+16 }
+17
+18 // H
+19
+20 P1(int *x, int **y)
+21 {
+22   int *r0;
+23   int r1;
+24
+25   rcu_read_lock();
+26   r0 = rcu_dereference(*y);
+27   r1 = READ_ONCE(*r0);
+28   rcu_read_unlock();
+29 }
+30
+31 // I
+32
+33 exists (* J *) (1:r0=x /\ (* K *) 1:r1=0) (* L *)
+
+In short, use C-language comments in the C code and Ocaml comments in
+the rest of the litmus test.
+
+On the other hand, if you prefer C-style comments everywhere, the
+C preprocessor is your friend.
+
+
+Asynchronous RCU Grace Periods
+------------------------------
+
+The following litmus test is derived from the example show in
+Documentation/litmus-tests/rcu/RCU+sync+free.litmus, but converted to
+emulate call_rcu():
+
+ 1 C RCU+sync+free
+ 2
+ 3 {
+ 4 int x = 1;
+ 5 int *y = &x;
+ 6 int z = 1;
+ 7 }
+ 8
+ 9 P0(int *x, int *z, int **y)
+10 {
+11   int *r0;
+12   int r1;
+13
+14   rcu_read_lock();
+15   r0 = rcu_dereference(*y);
+16   r1 = READ_ONCE(*r0);
+17   rcu_read_unlock();
+18 }
+19
+20 P1(int *z, int **y, int *c)
+21 {
+22   rcu_assign_pointer(*y, z);
+23   smp_store_release(*c, 1); // Emulate call_rcu().
+24 }
+25
+26 P2(int *x, int *z, int **y, int *c)
+27 {
+28   int r0;
+29
+30   r0 = smp_load_acquire(*c); // Note call_rcu() request.
+31   synchronize_rcu(); // Wait one grace period.
+32   WRITE_ONCE(*x, 0); // Emulate the RCU callback.
+33 }
+34
+35 filter (2:r0=1) (* Reject too-early starts. *)
+36 exists (0:r0=x /\ 0:r1=0)
+
+Lines 4-6 initialize a linked list headed by "y" that initially contains
+"x".  In addition, "z" is pre-initialized to prepare for P1(), which
+will replace "x" with "z" in this list.
+
+P0() on lines 9-18 enters an RCU read-side critical section, loads the
+list header "y" and dereferences it, leaving the node in "0:r0" and
+the node's value in "0:r1".
+
+P1() on lines 20-24 updates the list header to instead reference "z",
+then emulates call_rcu() by doing a release store into "c".
+
+P2() on lines 27-33 emulates the behind-the-scenes effect of doing a
+call_rcu().  Line 30 first does an acquire load from "c", then line 31
+waits for an RCU grace period to elapse, and finally line 32 emulates
+the RCU callback, which in turn emulates a call to kfree().
+
+Of course, it is possible for P2() to start too soon, so that the
+value of "2:r0" is zero rather than the required value of "1".
+The "filter" clause on line 35 handles this possibility, rejecting
+all executions in which "2:r0" is not equal to the value "1".
+
+
+Performance
+-----------
+
+LKMM's exploration of the full state-space can be extremely helpful,
+but it does not come for free.  The price is exponential computational
+complexity in terms of the number of processes, the average number
+of statements in each process, and the total number of stores in the
+litmus test.
+
+So it is best to start small and then work up.  Where possible, break
+your code down into small pieces each representing a core concurrency
+requirement.
+
+That said, herd7 is quite fast.  On an unprepossessing x86 laptop, it
+was able to analyze the following 10-process RCU litmus test in about
+six seconds.
+
+https://github.com/paulmckrcu/litmus/blob/master/auto/C-RW-R+RW-R+RW-G+RW-G+RW-G+RW-G+RW-R+RW-R+RW-R+RW-R.litmus
+
+One way to make herd7 run faster is to use the "-speedcheck true" option.
+This option prevents herd7 from generating all possible end states,
+instead causing it to focus solely on whether or not the "exists"
+clause can be satisfied.  With this option, herd7 evaluates the above
+litmus test in about 300 milliseconds, for more than an order of magnitude
+improvement in performance.
+
+Larger 16-process litmus tests that would normally consume 15 minutes
+of time complete in about 40 seconds with this option.  To be fair,
+you do get an extra 65,535 states when you leave off the "-speedcheck
+true" option.
+
+https://github.com/paulmckrcu/litmus/blob/master/auto/C-RW-R+RW-R+RW-G+RW-G+RW-G+RW-G+RW-R+RW-R+RW-R+RW-R+RW-G+RW-G+RW-G+RW-G+RW-R+RW-R.litmus
+
+Nevertheless, litmus-test analysis really is of exponential complexity,
+whether with or without "-speedcheck true".  Increasing by just three
+processes to a 19-process litmus test requires 2 hours and 40 minutes
+without, and about 8 minutes with "-speedcheck true".  Each of these
+results represent roughly an order of magnitude slowdown compared to the
+16-process litmus test.  Again, to be fair, the multi-hour run explores
+no fewer than 524,287 additional states compared to the shorter one.
+
+https://github.com/paulmckrcu/litmus/blob/master/auto/C-RW-R+RW-R+RW-G+RW-G+RW-G+RW-G+RW-R+RW-R+RW-R+RW-R+RW-R+RW-R+RW-G+RW-G+RW-G+RW-G+RW-R+RW-R+RW-R.litmus
+
+If you don't like command-line arguments, you can obtain a similar speedup
+by adding a "filter" clause with exactly the same expression as your
+"exists" clause.
+
+However, please note that seeing the full set of states can be extremely
+helpful when developing and debugging litmus tests.
+
+
+LIMITATIONS
+===========
+
+Limitations of the Linux-kernel memory model (LKMM) include:
+
+1.	Compiler optimizations are not accurately modeled.  Of course,
+	the use of READ_ONCE() and WRITE_ONCE() limits the compiler's
+	ability to optimize, but under some circumstances it is possible
+	for the compiler to undermine the memory model.  For more
+	information, see Documentation/explanation.txt (in particular,
+	the "THE PROGRAM ORDER RELATION: po AND po-loc" and "A WARNING"
+	sections).
+
+	Note that this limitation in turn limits LKMM's ability to
+	accurately model address, control, and data dependencies.
+	For example, if the compiler can deduce the value of some variable
+	carrying a dependency, then the compiler can break that dependency
+	by substituting a constant of that value.
+
+	Conversely, LKMM will sometimes overestimate the amount of
+	reordering compilers and CPUs can carry out, leading it to miss
+	some pretty obvious cases of ordering.  A simple example is:
+
+		r1 = READ_ONCE(x);
+		if (r1 == 0)
+			smp_mb();
+		WRITE_ONCE(y, 1);
+
+	The WRITE_ONCE() does not depend on the READ_ONCE(), and as a
+	result, LKMM does not claim ordering.  However, even though no
+	dependency is present, the WRITE_ONCE() will not be executed before
+	the READ_ONCE().  There are two reasons for this:
+
+                The presence of the smp_mb() in one of the branches
+                prevents the compiler from moving the WRITE_ONCE()
+                up before the "if" statement, since the compiler has
+                to assume that r1 will sometimes be 0 (but see the
+                comment below);
+
+                CPUs do not execute stores before po-earlier conditional
+                branches, even in cases where the store occurs after the
+                two arms of the branch have recombined.
+
+	It is clear that it is not dangerous in the slightest for LKMM to
+	make weaker guarantees than architectures.  In fact, it is
+	desirable, as it gives compilers room for making optimizations.
+	For instance, suppose that a 0 value in r1 would trigger undefined
+	behavior elsewhere.  Then a clever compiler might deduce that r1
+	can never be 0 in the if condition.  As a result, said clever
+	compiler might deem it safe to optimize away the smp_mb(),
+	eliminating the branch and any ordering an architecture would
+	guarantee otherwise.
+
+2.	Multiple access sizes for a single variable are not supported,
+	and neither are misaligned or partially overlapping accesses.
+
+3.	Exceptions and interrupts are not modeled.  In some cases,
+	this limitation can be overcome by modeling the interrupt or
+	exception with an additional process.
+
+4.	I/O such as MMIO or DMA is not supported.
+
+5.	Self-modifying code (such as that found in the kernel's
+	alternatives mechanism, function tracer, Berkeley Packet Filter
+	JIT compiler, and module loader) is not supported.
+
+6.	Complete modeling of all variants of atomic read-modify-write
+	operations, locking primitives, and RCU is not provided.
+	For example, call_rcu() and rcu_barrier() are not supported.
+	However, a substantial amount of support is provided for these
+	operations, as shown in the linux-kernel.def file.
+
+	Here are specific limitations:
+
+	a.	When rcu_assign_pointer() is passed NULL, the Linux
+		kernel provides no ordering, but LKMM models this
+		case as a store release.
+
+	b.	The "unless" RMW operations are not currently modeled:
+		atomic_long_add_unless(), atomic_inc_unless_negative(),
+		and atomic_dec_unless_positive().  These can be emulated
+		in litmus tests, for example, by using atomic_cmpxchg().
+
+		One exception of this limitation is atomic_add_unless(),
+		which is provided directly by herd7 (so no corresponding
+		definition in linux-kernel.def).  atomic_add_unless() is
+		modeled by herd7 therefore it can be used in litmus tests.
+
+	c.	The call_rcu() function is not modeled.  As was shown above,
+		it can be emulated in litmus tests by adding another
+		process that invokes synchronize_rcu() and the body of the
+		callback function, with (for example) a release-acquire
+		from the site of the emulated call_rcu() to the beginning
+		of the additional process.
+
+	d.	The rcu_barrier() function is not modeled.  It can be
+		emulated in litmus tests emulating call_rcu() via
+		(for example) a release-acquire from the end of each
+		additional call_rcu() process to the site of the
+		emulated rcu-barrier().
+
+	e.	Reader-writer locking is not modeled.  It can be
+		emulated in litmus tests using atomic read-modify-write
+		operations.
+
+The fragment of the C language supported by these litmus tests is quite
+limited and in some ways non-standard:
+
+1.	There is no automatic C-preprocessor pass.  You can of course
+	run it manually, if you choose.
+
+2.	There is no way to create functions other than the Pn() functions
+	that model the concurrent processes.
+
+3.	The Pn() functions' formal parameters must be pointers to the
+	global shared variables.  Nothing can be passed by value into
+	these functions.
+
+4.	The only functions that can be invoked are those built directly
+	into herd7 or that are defined in the linux-kernel.def file.
+
+5.	The "switch", "do", "for", "while", and "goto" C statements are
+	not supported.	The "switch" statement can be emulated by the
+	"if" statement.  The "do", "for", and "while" statements can
+	often be emulated by manually unrolling the loop, or perhaps by
+	enlisting the aid of the C preprocessor to minimize the resulting
+	code duplication.  Some uses of "goto" can be emulated by "if",
+	and some others by unrolling.
+
+6.	Although you can use a wide variety of types in litmus-test
+	variable declarations, and especially in global-variable
+	declarations, the "herd7" tool understands only int and
+	pointer types.	There is no support for floating-point types,
+	enumerations, characters, strings, arrays, or structures.
+
+7.	Parsing of variable declarations is very loose, with almost no
+	type checking.
+
+8.	Initializers differ from their C-language counterparts.
+	For example, when an initializer contains the name of a shared
+	variable, that name denotes a pointer to that variable, not
+	the current value of that variable.  For example, "int x = y"
+	is interpreted the way "int x = &y" would be in C.
+
+9.	Dynamic memory allocation is not supported, although this can
+	be worked around in some cases by supplying multiple statically
+	allocated variables.
+
+Some of these limitations may be overcome in the future, but others are
+more likely to be addressed by incorporating the Linux-kernel memory model
+into other tools.
+
+Finally, please note that LKMM is subject to change as hardware, use cases,
+and compilers evolve.
diff --git a/tools/memory-model/Documentation/locking.txt b/tools/memory-model/Documentation/locking.txt
new file mode 100644
index 0000000000..65c898c64a
--- /dev/null
+++ b/tools/memory-model/Documentation/locking.txt
@@ -0,0 +1,298 @@
+Locking
+=======
+
+Locking is well-known and the common use cases are straightforward: Any
+CPU holding a given lock sees any changes previously seen or made by any
+CPU before it previously released that same lock.  This last sentence
+is the only part of this document that most developers will need to read.
+
+However, developers who would like to also access lock-protected shared
+variables outside of their corresponding locks should continue reading.
+
+
+Locking and Prior Accesses
+--------------------------
+
+The basic rule of locking is worth repeating:
+
+	Any CPU holding a given lock sees any changes previously seen
+	or made by any CPU before it previously released that same lock.
+
+Note that this statement is a bit stronger than "Any CPU holding a
+given lock sees all changes made by any CPU during the time that CPU was
+previously holding this same lock".  For example, consider the following
+pair of code fragments:
+
+	/* See MP+polocks.litmus. */
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		spin_lock(&mylock);
+		WRITE_ONCE(y, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		r0 = READ_ONCE(y);
+		spin_unlock(&mylock);
+		r1 = READ_ONCE(x);
+	}
+
+The basic rule guarantees that if CPU0() acquires mylock before CPU1(),
+then both r0 and r1 must be set to the value 1.  This also has the
+consequence that if the final value of r0 is equal to 1, then the final
+value of r1 must also be equal to 1.  In contrast, the weaker rule would
+say nothing about the final value of r1.
+
+
+Locking and Subsequent Accesses
+-------------------------------
+
+The converse to the basic rule also holds:  Any CPU holding a given
+lock will not see any changes that will be made by any CPU after it
+subsequently acquires this same lock.  This converse statement is
+illustrated by the following litmus test:
+
+	/* See MP+porevlocks.litmus. */
+	void CPU0(void)
+	{
+		r0 = READ_ONCE(y);
+		spin_lock(&mylock);
+		r1 = READ_ONCE(x);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		WRITE_ONCE(x, 1);
+		spin_unlock(&mylock);
+		WRITE_ONCE(y, 1);
+	}
+
+This converse to the basic rule guarantees that if CPU0() acquires
+mylock before CPU1(), then both r0 and r1 must be set to the value 0.
+This also has the consequence that if the final value of r1 is equal
+to 0, then the final value of r0 must also be equal to 0.  In contrast,
+the weaker rule would say nothing about the final value of r0.
+
+These examples show only a single pair of CPUs, but the effects of the
+locking basic rule extend across multiple acquisitions of a given lock
+across multiple CPUs.
+
+
+Double-Checked Locking
+----------------------
+
+It is well known that more than just a lock is required to make
+double-checked locking work correctly,  This litmus test illustrates
+one incorrect approach:
+
+	/* See Documentation/litmus-tests/locking/DCL-broken.litmus. */
+	void CPU0(void)
+	{
+		r0 = READ_ONCE(flag);
+		if (r0 == 0) {
+			spin_lock(&lck);
+			r1 = READ_ONCE(flag);
+			if (r1 == 0) {
+				WRITE_ONCE(data, 1);
+				WRITE_ONCE(flag, 1);
+			}
+			spin_unlock(&lck);
+		}
+		r2 = READ_ONCE(data);
+	}
+	/* CPU1() is the exactly the same as CPU0(). */
+
+There are two problems.  First, there is no ordering between the first
+READ_ONCE() of "flag" and the READ_ONCE() of "data".  Second, there is
+no ordering between the two WRITE_ONCE() calls.  It should therefore be
+no surprise that "r2" can be zero, and a quick herd7 run confirms this.
+
+One way to fix this is to use smp_load_acquire() and smp_store_release()
+as shown in this corrected version:
+
+	/* See Documentation/litmus-tests/locking/DCL-fixed.litmus. */
+	void CPU0(void)
+	{
+		r0 = smp_load_acquire(&flag);
+		if (r0 == 0) {
+			spin_lock(&lck);
+			r1 = READ_ONCE(flag);
+			if (r1 == 0) {
+				WRITE_ONCE(data, 1);
+				smp_store_release(&flag, 1);
+			}
+			spin_unlock(&lck);
+		}
+		r2 = READ_ONCE(data);
+	}
+	/* CPU1() is the exactly the same as CPU0(). */
+
+The smp_load_acquire() guarantees that its load from "flags" will
+be ordered before the READ_ONCE() from data, thus solving the first
+problem.  The smp_store_release() guarantees that its store will be
+ordered after the WRITE_ONCE() to "data", solving the second problem.
+The smp_store_release() pairs with the smp_load_acquire(), thus ensuring
+that the ordering provided by each actually takes effect.  Again, a
+quick herd7 run confirms this.
+
+In short, if you access a lock-protected variable without holding the
+corresponding lock, you will need to provide additional ordering, in
+this case, via the smp_load_acquire() and the smp_store_release().
+
+
+Ordering Provided by a Lock to CPUs Not Holding That Lock
+---------------------------------------------------------
+
+It is not necessarily the case that accesses ordered by locking will be
+seen as ordered by CPUs not holding that lock.  Consider this example:
+
+	/* See Z6.0+pooncelock+pooncelock+pombonce.litmus. */
+	void CPU0(void)
+	{
+		spin_lock(&mylock);
+		WRITE_ONCE(x, 1);
+		WRITE_ONCE(y, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		r0 = READ_ONCE(y);
+		WRITE_ONCE(z, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU2(void)
+	{
+		WRITE_ONCE(z, 2);
+		smp_mb();
+		r1 = READ_ONCE(x);
+	}
+
+Counter-intuitive though it might be, it is quite possible to have
+the final value of r0 be 1, the final value of z be 2, and the final
+value of r1 be 0.  The reason for this surprising outcome is that CPU2()
+never acquired the lock, and thus did not fully benefit from the lock's
+ordering properties.
+
+Ordering can be extended to CPUs not holding the lock by careful use
+of smp_mb__after_spinlock():
+
+	/* See Z6.0+pooncelock+poonceLock+pombonce.litmus. */
+	void CPU0(void)
+	{
+		spin_lock(&mylock);
+		WRITE_ONCE(x, 1);
+		WRITE_ONCE(y, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		smp_mb__after_spinlock();
+		r0 = READ_ONCE(y);
+		WRITE_ONCE(z, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU2(void)
+	{
+		WRITE_ONCE(z, 2);
+		smp_mb();
+		r1 = READ_ONCE(x);
+	}
+
+This addition of smp_mb__after_spinlock() strengthens the lock
+acquisition sufficiently to rule out the counter-intuitive outcome.
+In other words, the addition of the smp_mb__after_spinlock() prohibits
+the counter-intuitive result where the final value of r0 is 1, the final
+value of z is 2, and the final value of r1 is 0.
+
+
+No Roach-Motel Locking!
+-----------------------
+
+This example requires familiarity with the herd7 "filter" clause, so
+please read up on that topic in litmus-tests.txt.
+
+It is tempting to allow memory-reference instructions to be pulled
+into a critical section, but this cannot be allowed in the general case.
+For example, consider a spin loop preceding a lock-based critical section.
+Now, herd7 does not model spin loops, but we can emulate one with two
+loads, with a "filter" clause to constrain the first to return the
+initial value and the second to return the updated value, as shown below:
+
+	/* See Documentation/litmus-tests/locking/RM-fixed.litmus. */
+	void CPU0(void)
+	{
+		spin_lock(&lck);
+		r2 = atomic_inc_return(&y);
+		WRITE_ONCE(x, 1);
+		spin_unlock(&lck);
+	}
+
+	void CPU1(void)
+	{
+		r0 = READ_ONCE(x);
+		r1 = READ_ONCE(x);
+		spin_lock(&lck);
+		r2 = atomic_inc_return(&y);
+		spin_unlock(&lck);
+	}
+
+	filter (1:r0=0 /\ 1:r1=1)
+	exists (1:r2=1)
+
+The variable "x" is the control variable for the emulated spin loop.
+CPU0() sets it to "1" while holding the lock, and CPU1() emulates the
+spin loop by reading it twice, first into "1:r0" (which should get the
+initial value "0") and then into "1:r1" (which should get the updated
+value "1").
+
+The "filter" clause takes this into account, constraining "1:r0" to
+equal "0" and "1:r1" to equal 1.
+
+Then the "exists" clause checks to see if CPU1() acquired its lock first,
+which should not happen given the filter clause because CPU0() updates
+"x" while holding the lock.  And herd7 confirms this.
+
+But suppose that the compiler was permitted to reorder the spin loop
+into CPU1()'s critical section, like this:
+
+	/* See Documentation/litmus-tests/locking/RM-broken.litmus. */
+	void CPU0(void)
+	{
+		int r2;
+
+		spin_lock(&lck);
+		r2 = atomic_inc_return(&y);
+		WRITE_ONCE(x, 1);
+		spin_unlock(&lck);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&lck);
+		r0 = READ_ONCE(x);
+		r1 = READ_ONCE(x);
+		r2 = atomic_inc_return(&y);
+		spin_unlock(&lck);
+	}
+
+	filter (1:r0=0 /\ 1:r1=1)
+	exists (1:r2=1)
+
+If "1:r0" is equal to "0", "1:r1" can never equal "1" because CPU0()
+cannot update "x" while CPU1() holds the lock.  And herd7 confirms this,
+showing zero executions matching the "filter" criteria.
+
+And this is why Linux-kernel lock and unlock primitives must prevent
+code from entering critical sections.  It is not sufficient to only
+prevent code from leaving them.
diff --git a/tools/memory-model/Documentation/ordering.txt b/tools/memory-model/Documentation/ordering.txt
new file mode 100644
index 0000000000..9b0949d3f5
--- /dev/null
+++ b/tools/memory-model/Documentation/ordering.txt
@@ -0,0 +1,556 @@
+This document gives an overview of the categories of memory-ordering
+operations provided by the Linux-kernel memory model (LKMM).
+
+
+Categories of Ordering
+======================
+
+This section lists LKMM's three top-level categories of memory-ordering
+operations in decreasing order of strength:
+
+1.	Barriers (also known as "fences").  A barrier orders some or
+	all of the CPU's prior operations against some or all of its
+	subsequent operations.
+
+2.	Ordered memory accesses.  These operations order themselves
+	against some or all of the CPU's prior accesses or some or all
+	of the CPU's subsequent accesses, depending on the subcategory
+	of the operation.
+
+3.	Unordered accesses, as the name indicates, have no ordering
+	properties except to the extent that they interact with an
+	operation in the previous categories.  This being the real world,
+	some of these "unordered" operations provide limited ordering
+	in some special situations.
+
+Each of the above categories is described in more detail by one of the
+following sections.
+
+
+Barriers
+========
+
+Each of the following categories of barriers is described in its own
+subsection below:
+
+a.	Full memory barriers.
+
+b.	Read-modify-write (RMW) ordering augmentation barriers.
+
+c.	Write memory barrier.
+
+d.	Read memory barrier.
+
+e.	Compiler barrier.
+
+Note well that many of these primitives generate absolutely no code
+in kernels built with CONFIG_SMP=n.  Therefore, if you are writing
+a device driver, which must correctly order accesses to a physical
+device even in kernels built with CONFIG_SMP=n, please use the
+ordering primitives provided for that purpose.  For example, instead of
+smp_mb(), use mb().  See the "Linux Kernel Device Drivers" book or the
+https://lwn.net/Articles/698014/ article for more information.
+
+
+Full Memory Barriers
+--------------------
+
+The Linux-kernel primitives that provide full ordering include:
+
+o	The smp_mb() full memory barrier.
+
+o	Value-returning RMW atomic operations whose names do not end in
+	_acquire, _release, or _relaxed.
+
+o	RCU's grace-period primitives.
+
+First, the smp_mb() full memory barrier orders all of the CPU's prior
+accesses against all subsequent accesses from the viewpoint of all CPUs.
+In other words, all CPUs will agree that any earlier action taken
+by that CPU happened before any later action taken by that same CPU.
+For example, consider the following:
+
+	WRITE_ONCE(x, 1);
+	smp_mb(); // Order store to x before load from y.
+	r1 = READ_ONCE(y);
+
+All CPUs will agree that the store to "x" happened before the load
+from "y", as indicated by the comment.  And yes, please comment your
+memory-ordering primitives.  It is surprisingly hard to remember their
+purpose after even a few months.
+
+Second, some RMW atomic operations provide full ordering.  These
+operations include value-returning RMW atomic operations (that is, those
+with non-void return types) whose names do not end in _acquire, _release,
+or _relaxed.  Examples include atomic_add_return(), atomic_dec_and_test(),
+cmpxchg(), and xchg().  Note that conditional RMW atomic operations such
+as cmpxchg() are only guaranteed to provide ordering when they succeed.
+When RMW atomic operations provide full ordering, they partition the
+CPU's accesses into three groups:
+
+1.	All code that executed prior to the RMW atomic operation.
+
+2.	The RMW atomic operation itself.
+
+3.	All code that executed after the RMW atomic operation.
+
+All CPUs will agree that any operation in a given partition happened
+before any operation in a higher-numbered partition.
+
+In contrast, non-value-returning RMW atomic operations (that is, those
+with void return types) do not guarantee any ordering whatsoever.  Nor do
+value-returning RMW atomic operations whose names end in _relaxed.
+Examples of the former include atomic_inc() and atomic_dec(),
+while examples of the latter include atomic_cmpxchg_relaxed() and
+atomic_xchg_relaxed().  Similarly, value-returning non-RMW atomic
+operations such as atomic_read() do not guarantee full ordering, and
+are covered in the later section on unordered operations.
+
+Value-returning RMW atomic operations whose names end in _acquire or
+_release provide limited ordering, and will be described later in this
+document.
+
+Finally, RCU's grace-period primitives provide full ordering.  These
+primitives include synchronize_rcu(), synchronize_rcu_expedited(),
+synchronize_srcu() and so on.  However, these primitives have orders
+of magnitude greater overhead than smp_mb(), atomic_xchg(), and so on.
+Furthermore, RCU's grace-period primitives can only be invoked in
+sleepable contexts.  Therefore, RCU's grace-period primitives are
+typically instead used to provide ordering against RCU read-side critical
+sections, as documented in their comment headers.  But of course if you
+need a synchronize_rcu() to interact with readers, it costs you nothing
+to also rely on its additional full-memory-barrier semantics.  Just please
+carefully comment this, otherwise your future self will hate you.
+
+
+RMW Ordering Augmentation Barriers
+----------------------------------
+
+As noted in the previous section, non-value-returning RMW operations
+such as atomic_inc() and atomic_dec() guarantee no ordering whatsoever.
+Nevertheless, a number of popular CPU families, including x86, provide
+full ordering for these primitives.  One way to obtain full ordering on
+all architectures is to add a call to smp_mb():
+
+	WRITE_ONCE(x, 1);
+	atomic_inc(&my_counter);
+	smp_mb(); // Inefficient on x86!!!
+	r1 = READ_ONCE(y);
+
+This works, but the added smp_mb() adds needless overhead for
+x86, on which atomic_inc() provides full ordering all by itself.
+The smp_mb__after_atomic() primitive can be used instead:
+
+	WRITE_ONCE(x, 1);
+	atomic_inc(&my_counter);
+	smp_mb__after_atomic(); // Order store to x before load from y.
+	r1 = READ_ONCE(y);
+
+The smp_mb__after_atomic() primitive emits code only on CPUs whose
+atomic_inc() implementations do not guarantee full ordering, thus
+incurring no unnecessary overhead on x86.  There are a number of
+variations on the smp_mb__*() theme:
+
+o	smp_mb__before_atomic(), which provides full ordering prior
+	to an unordered RMW atomic operation.
+
+o	smp_mb__after_atomic(), which, as shown above, provides full
+	ordering subsequent to an unordered RMW atomic operation.
+
+o	smp_mb__after_spinlock(), which provides full ordering subsequent
+	to a successful spinlock acquisition.  Note that spin_lock() is
+	always successful but spin_trylock() might not be.
+
+o	smp_mb__after_srcu_read_unlock(), which provides full ordering
+	subsequent to an srcu_read_unlock().
+
+It is bad practice to place code between the smp__*() primitive and the
+operation whose ordering that it is augmenting.  The reason is that the
+ordering of this intervening code will differ from one CPU architecture
+to another.
+
+
+Write Memory Barrier
+--------------------
+
+The Linux kernel's write memory barrier is smp_wmb().  If a CPU executes
+the following code:
+
+	WRITE_ONCE(x, 1);
+	smp_wmb();
+	WRITE_ONCE(y, 1);
+
+Then any given CPU will see the write to "x" has having happened before
+the write to "y".  However, you are usually better off using a release
+store, as described in the "Release Operations" section below.
+
+Note that smp_wmb() might fail to provide ordering for unmarked C-language
+stores because profile-driven optimization could determine that the
+value being overwritten is almost always equal to the new value.  Such a
+compiler might then reasonably decide to transform "x = 1" and "y = 1"
+as follows:
+
+	if (x != 1)
+		x = 1;
+	smp_wmb(); // BUG: does not order the reads!!!
+	if (y != 1)
+		y = 1;
+
+Therefore, if you need to use smp_wmb() with unmarked C-language writes,
+you will need to make sure that none of the compilers used to build
+the Linux kernel carry out this sort of transformation, both now and in
+the future.
+
+
+Read Memory Barrier
+-------------------
+
+The Linux kernel's read memory barrier is smp_rmb().  If a CPU executes
+the following code:
+
+	r0 = READ_ONCE(y);
+	smp_rmb();
+	r1 = READ_ONCE(x);
+
+Then any given CPU will see the read from "y" as having preceded the read from
+"x".  However, you are usually better off using an acquire load, as described
+in the "Acquire Operations" section below.
+
+Compiler Barrier
+----------------
+
+The Linux kernel's compiler barrier is barrier().  This primitive
+prohibits compiler code-motion optimizations that might move memory
+references across the point in the code containing the barrier(), but
+does not constrain hardware memory ordering.  For example, this can be
+used to prevent to compiler from moving code across an infinite loop:
+
+	WRITE_ONCE(x, 1);
+	while (dontstop)
+		barrier();
+	r1 = READ_ONCE(y);
+
+Without the barrier(), the compiler would be within its rights to move the
+WRITE_ONCE() to follow the loop.  This code motion could be problematic
+in the case where an interrupt handler terminates the loop.  Another way
+to handle this is to use READ_ONCE() for the load of "dontstop".
+
+Note that the barriers discussed previously use barrier() or its low-level
+equivalent in their implementations.
+
+
+Ordered Memory Accesses
+=======================
+
+The Linux kernel provides a wide variety of ordered memory accesses:
+
+a.	Release operations.
+
+b.	Acquire operations.
+
+c.	RCU read-side ordering.
+
+d.	Control dependencies.
+
+Each of the above categories has its own section below.
+
+
+Release Operations
+------------------
+
+Release operations include smp_store_release(), atomic_set_release(),
+rcu_assign_pointer(), and value-returning RMW operations whose names
+end in _release.  These operations order their own store against all
+of the CPU's prior memory accesses.  Release operations often provide
+improved readability and performance compared to explicit barriers.
+For example, use of smp_store_release() saves a line compared to the
+smp_wmb() example above:
+
+	WRITE_ONCE(x, 1);
+	smp_store_release(&y, 1);
+
+More important, smp_store_release() makes it easier to connect up the
+different pieces of the concurrent algorithm.  The variable stored to
+by the smp_store_release(), in this case "y", will normally be used in
+an acquire operation in other parts of the concurrent algorithm.
+
+To see the performance advantages, suppose that the above example read
+from "x" instead of writing to it.  Then an smp_wmb() could not guarantee
+ordering, and an smp_mb() would be needed instead:
+
+	r1 = READ_ONCE(x);
+	smp_mb();
+	WRITE_ONCE(y, 1);
+
+But smp_mb() often incurs much higher overhead than does
+smp_store_release(), which still provides the needed ordering of "x"
+against "y".  On x86, the version using smp_store_release() might compile
+to a simple load instruction followed by a simple store instruction.
+In contrast, the smp_mb() compiles to an expensive instruction that
+provides the needed ordering.
+
+There is a wide variety of release operations:
+
+o	Store operations, including not only the aforementioned
+	smp_store_release(), but also atomic_set_release(), and
+	atomic_long_set_release().
+
+o	RCU's rcu_assign_pointer() operation.  This is the same as
+	smp_store_release() except that: (1) It takes the pointer to
+	be assigned to instead of a pointer to that pointer, (2) It
+	is intended to be used in conjunction with rcu_dereference()
+	and similar rather than smp_load_acquire(), and (3) It checks
+	for an RCU-protected pointer in "sparse" runs.
+
+o	Value-returning RMW operations whose names end in _release,
+	such as atomic_fetch_add_release() and cmpxchg_release().
+	Note that release ordering is guaranteed only against the
+	memory-store portion of the RMW operation, and not against the
+	memory-load portion.  Note also that conditional operations such
+	as cmpxchg_release() are only guaranteed to provide ordering
+	when they succeed.
+
+As mentioned earlier, release operations are often paired with acquire
+operations, which are the subject of the next section.
+
+
+Acquire Operations
+------------------
+
+Acquire operations include smp_load_acquire(), atomic_read_acquire(),
+and value-returning RMW operations whose names end in _acquire.   These
+operations order their own load against all of the CPU's subsequent
+memory accesses.  Acquire operations often provide improved performance
+and readability compared to explicit barriers.  For example, use of
+smp_load_acquire() saves a line compared to the smp_rmb() example above:
+
+	r0 = smp_load_acquire(&y);
+	r1 = READ_ONCE(x);
+
+As with smp_store_release(), this also makes it easier to connect
+the different pieces of the concurrent algorithm by looking for the
+smp_store_release() that stores to "y".  In addition, smp_load_acquire()
+improves upon smp_rmb() by ordering against subsequent stores as well
+as against subsequent loads.
+
+There are a couple of categories of acquire operations:
+
+o	Load operations, including not only the aforementioned
+	smp_load_acquire(), but also atomic_read_acquire(), and
+	atomic64_read_acquire().
+
+o	Value-returning RMW operations whose names end in _acquire,
+	such as atomic_xchg_acquire() and atomic_cmpxchg_acquire().
+	Note that acquire ordering is guaranteed only against the
+	memory-load portion of the RMW operation, and not against the
+	memory-store portion.  Note also that conditional operations
+	such as atomic_cmpxchg_acquire() are only guaranteed to provide
+	ordering when they succeed.
+
+Symmetry being what it is, acquire operations are often paired with the
+release operations covered earlier.  For example, consider the following
+example, where task0() and task1() execute concurrently:
+
+	void task0(void)
+	{
+		WRITE_ONCE(x, 1);
+		smp_store_release(&y, 1);
+	}
+
+	void task1(void)
+	{
+		r0 = smp_load_acquire(&y);
+		r1 = READ_ONCE(x);
+	}
+
+If "x" and "y" are both initially zero, then either r0's final value
+will be zero or r1's final value will be one, thus providing the required
+ordering.
+
+
+RCU Read-Side Ordering
+----------------------
+
+This category includes read-side markers such as rcu_read_lock()
+and rcu_read_unlock() as well as pointer-traversal primitives such as
+rcu_dereference() and srcu_dereference().
+
+Compared to locking primitives and RMW atomic operations, markers
+for RCU read-side critical sections incur very low overhead because
+they interact only with the corresponding grace-period primitives.
+For example, the rcu_read_lock() and rcu_read_unlock() markers interact
+with synchronize_rcu(), synchronize_rcu_expedited(), and call_rcu().
+The way this works is that if a given call to synchronize_rcu() cannot
+prove that it started before a given call to rcu_read_lock(), then
+that synchronize_rcu() must block until the matching rcu_read_unlock()
+is reached.  For more information, please see the synchronize_rcu()
+docbook header comment and the material in Documentation/RCU.
+
+RCU's pointer-traversal primitives, including rcu_dereference() and
+srcu_dereference(), order their load (which must be a pointer) against any
+of the CPU's subsequent memory accesses whose address has been calculated
+from the value loaded.  There is said to be an *address dependency*
+from the value returned by the rcu_dereference() or srcu_dereference()
+to that subsequent memory access.
+
+A call to rcu_dereference() for a given RCU-protected pointer is
+usually paired with a call to a call to rcu_assign_pointer() for that
+same pointer in much the same way that a call to smp_load_acquire() is
+paired with a call to smp_store_release().  Calls to rcu_dereference()
+and rcu_assign_pointer are often buried in other APIs, for example,
+the RCU list API members defined in include/linux/rculist.h.  For more
+information, please see the docbook headers in that file, the most
+recent LWN article on the RCU API (https://lwn.net/Articles/777036/),
+and of course the material in Documentation/RCU.
+
+If the pointer value is manipulated between the rcu_dereference()
+that returned it and a later dereference(), please read
+Documentation/RCU/rcu_dereference.rst.  It can also be quite helpful to
+review uses in the Linux kernel.
+
+
+Control Dependencies
+--------------------
+
+A control dependency extends from a marked load (READ_ONCE() or stronger)
+through an "if" condition to a marked store (WRITE_ONCE() or stronger)
+that is executed only by one of the legs of that "if" statement.
+Control dependencies are so named because they are mediated by
+control-flow instructions such as comparisons and conditional branches.
+
+In short, you can use a control dependency to enforce ordering between
+an READ_ONCE() and a WRITE_ONCE() when there is an "if" condition
+between them.  The canonical example is as follows:
+
+	q = READ_ONCE(a);
+	if (q)
+		WRITE_ONCE(b, 1);
+
+In this case, all CPUs would see the read from "a" as happening before
+the write to "b".
+
+However, control dependencies are easily destroyed by compiler
+optimizations, so any use of control dependencies must take into account
+all of the compilers used to build the Linux kernel.  Please see the
+"control-dependencies.txt" file for more information.
+
+
+Unordered Accesses
+==================
+
+Each of these two categories of unordered accesses has a section below:
+
+a.	Unordered marked operations.
+
+b.	Unmarked C-language accesses.
+
+
+Unordered Marked Operations
+---------------------------
+
+Unordered operations to different variables are just that, unordered.
+However, if a group of CPUs apply these operations to a single variable,
+all the CPUs will agree on the operation order.  Of course, the ordering
+of unordered marked accesses can also be constrained using the mechanisms
+described earlier in this document.
+
+These operations come in three categories:
+
+o	Marked writes, such as WRITE_ONCE() and atomic_set().  These
+	primitives required the compiler to emit the corresponding store
+	instructions in the expected execution order, thus suppressing
+	a number of destructive optimizations.	However, they provide no
+	hardware ordering guarantees, and in fact many CPUs will happily
+	reorder marked writes with each other or with other unordered
+	operations, unless these operations are to the same variable.
+
+o	Marked reads, such as READ_ONCE() and atomic_read().  These
+	primitives required the compiler to emit the corresponding load
+	instructions in the expected execution order, thus suppressing
+	a number of destructive optimizations.	However, they provide no
+	hardware ordering guarantees, and in fact many CPUs will happily
+	reorder marked reads with each other or with other unordered
+	operations, unless these operations are to the same variable.
+
+o	Unordered RMW atomic operations.  These are non-value-returning
+	RMW atomic operations whose names do not end in _acquire or
+	_release, and also value-returning RMW operations whose names
+	end in _relaxed.  Examples include atomic_add(), atomic_or(),
+	and atomic64_fetch_xor_relaxed().  These operations do carry
+	out the specified RMW operation atomically, for example, five
+	concurrent atomic_inc() operations applied to a given variable
+	will reliably increase the value of that variable by five.
+	However, many CPUs will happily reorder these operations with
+	each other or with other unordered operations.
+
+	This category of operations can be efficiently ordered using
+	smp_mb__before_atomic() and smp_mb__after_atomic(), as was
+	discussed in the "RMW Ordering Augmentation Barriers" section.
+
+In short, these operations can be freely reordered unless they are all
+operating on a single variable or unless they are constrained by one of
+the operations called out earlier in this document.
+
+
+Unmarked C-Language Accesses
+----------------------------
+
+Unmarked C-language accesses are normal variable accesses to normal
+variables, that is, to variables that are not "volatile" and are not
+C11 atomic variables.  These operations provide no ordering guarantees,
+and further do not guarantee "atomic" access.  For example, the compiler
+might (and sometimes does) split a plain C-language store into multiple
+smaller stores.  A load from that same variable running on some other
+CPU while such a store is executing might see a value that is a mashup
+of the old value and the new value.
+
+Unmarked C-language accesses are unordered, and are also subject to
+any number of compiler optimizations, many of which can break your
+concurrent code.  It is possible to used unmarked C-language accesses for
+shared variables that are subject to concurrent access, but great care
+is required on an ongoing basis.  The compiler-constraining barrier()
+primitive can be helpful, as can the various ordering primitives discussed
+in this document.  It nevertheless bears repeating that use of unmarked
+C-language accesses requires careful attention to not just your code,
+but to all the compilers that might be used to build it.  Such compilers
+might replace a series of loads with a single load, and might replace
+a series of stores with a single store.  Some compilers will even split
+a single store into multiple smaller stores.
+
+But there are some ways of using unmarked C-language accesses for shared
+variables without such worries:
+
+o	Guard all accesses to a given variable by a particular lock,
+	so that there are never concurrent conflicting accesses to
+	that variable.	(There are "conflicting accesses" when
+	(1) at least one of the concurrent accesses to a variable is an
+	unmarked C-language access and (2) when at least one of those
+	accesses is a write, whether marked or not.)
+
+o	As above, but using other synchronization primitives such
+	as reader-writer locks or sequence locks.
+
+o	Use locking or other means to ensure that all concurrent accesses
+	to a given variable are reads.
+
+o	Restrict use of a given variable to statistics or heuristics
+	where the occasional bogus value can be tolerated.
+
+o	Declare the accessed variables as C11 atomics.
+	https://lwn.net/Articles/691128/
+
+o	Declare the accessed variables as "volatile".
+
+If you need to live more dangerously, please do take the time to
+understand the compilers.  One place to start is these two LWN
+articles:
+
+Who's afraid of a big bad optimizing compiler?
+	https://lwn.net/Articles/793253
+Calibrating your fear of big bad optimizing compilers
+	https://lwn.net/Articles/799218
+
+Used properly, unmarked C-language accesses can reduce overhead on
+fastpaths.  However, the price is great care and continual attention
+to your compiler as new versions come out and as new optimizations
+are enabled.
diff --git a/tools/memory-model/Documentation/recipes.txt b/tools/memory-model/Documentation/recipes.txt
new file mode 100644
index 0000000000..03f58b11c2
--- /dev/null
+++ b/tools/memory-model/Documentation/recipes.txt
@@ -0,0 +1,570 @@
+This document provides "recipes", that is, litmus tests for commonly
+occurring situations, as well as a few that illustrate subtly broken but
+attractive nuisances.  Many of these recipes include example code from
+v5.7 of the Linux kernel.
+
+The first section covers simple special cases, the second section
+takes off the training wheels to cover more involved examples,
+and the third section provides a few rules of thumb.
+
+
+Simple special cases
+====================
+
+This section presents two simple special cases, the first being where
+there is only one CPU or only one memory location is accessed, and the
+second being use of that old concurrency workhorse, locking.
+
+
+Single CPU or single memory location
+------------------------------------
+
+If there is only one CPU on the one hand or only one variable
+on the other, the code will execute in order.  There are (as
+usual) some things to be careful of:
+
+1.	Some aspects of the C language are unordered.  For example,
+	in the expression "f(x) + g(y)", the order in which f and g are
+	called is not defined; the object code is allowed to use either
+	order or even to interleave the computations.
+
+2.	Compilers are permitted to use the "as-if" rule.  That is, a
+	compiler can emit whatever code it likes for normal accesses,
+	as long as the results of a single-threaded execution appear
+	just as if the compiler had followed all the relevant rules.
+	To see this, compile with a high level of optimization and run
+	the debugger on the resulting binary.
+
+3.	If there is only one variable but multiple CPUs, that variable
+	must be properly aligned and all accesses to that variable must
+	be full sized.	Variables that straddle cachelines or pages void
+	your full-ordering warranty, as do undersized accesses that load
+	from or store to only part of the variable.
+
+4.	If there are multiple CPUs, accesses to shared variables should
+	use READ_ONCE() and WRITE_ONCE() or stronger to prevent load/store
+	tearing, load/store fusing, and invented loads and stores.
+	There are exceptions to this rule, including:
+
+	i.	When there is no possibility of a given shared variable
+		being updated by some other CPU, for example, while
+		holding the update-side lock, reads from that variable
+		need not use READ_ONCE().
+
+	ii.	When there is no possibility of a given shared variable
+		being either read or updated by other CPUs, for example,
+		when running during early boot, reads from that variable
+		need not use READ_ONCE() and writes to that variable
+		need not use WRITE_ONCE().
+
+
+Locking
+-------
+
+Locking is well-known and straightforward, at least if you don't think
+about it too hard.  And the basic rule is indeed quite simple: Any CPU that
+has acquired a given lock sees any changes previously seen or made by any
+CPU before it released that same lock.  Note that this statement is a bit
+stronger than "Any CPU holding a given lock sees all changes made by any
+CPU during the time that CPU was holding this same lock".  For example,
+consider the following pair of code fragments:
+
+	/* See MP+polocks.litmus. */
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		spin_lock(&mylock);
+		WRITE_ONCE(y, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		r0 = READ_ONCE(y);
+		spin_unlock(&mylock);
+		r1 = READ_ONCE(x);
+	}
+
+The basic rule guarantees that if CPU0() acquires mylock before CPU1(),
+then both r0 and r1 must be set to the value 1.  This also has the
+consequence that if the final value of r0 is equal to 1, then the final
+value of r1 must also be equal to 1.  In contrast, the weaker rule would
+say nothing about the final value of r1.
+
+The converse to the basic rule also holds, as illustrated by the
+following litmus test:
+
+	/* See MP+porevlocks.litmus. */
+	void CPU0(void)
+	{
+		r0 = READ_ONCE(y);
+		spin_lock(&mylock);
+		r1 = READ_ONCE(x);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		WRITE_ONCE(x, 1);
+		spin_unlock(&mylock);
+		WRITE_ONCE(y, 1);
+	}
+
+This converse to the basic rule guarantees that if CPU0() acquires
+mylock before CPU1(), then both r0 and r1 must be set to the value 0.
+This also has the consequence that if the final value of r1 is equal
+to 0, then the final value of r0 must also be equal to 0.  In contrast,
+the weaker rule would say nothing about the final value of r0.
+
+These examples show only a single pair of CPUs, but the effects of the
+locking basic rule extend across multiple acquisitions of a given lock
+across multiple CPUs.
+
+However, it is not necessarily the case that accesses ordered by
+locking will be seen as ordered by CPUs not holding that lock.
+Consider this example:
+
+	/* See Z6.0+pooncelock+pooncelock+pombonce.litmus. */
+	void CPU0(void)
+	{
+		spin_lock(&mylock);
+		WRITE_ONCE(x, 1);
+		WRITE_ONCE(y, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		r0 = READ_ONCE(y);
+		WRITE_ONCE(z, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU2(void)
+	{
+		WRITE_ONCE(z, 2);
+		smp_mb();
+		r1 = READ_ONCE(x);
+	}
+
+Counter-intuitive though it might be, it is quite possible to have
+the final value of r0 be 1, the final value of z be 2, and the final
+value of r1 be 0.  The reason for this surprising outcome is that
+CPU2() never acquired the lock, and thus did not benefit from the
+lock's ordering properties.
+
+Ordering can be extended to CPUs not holding the lock by careful use
+of smp_mb__after_spinlock():
+
+	/* See Z6.0+pooncelock+poonceLock+pombonce.litmus. */
+	void CPU0(void)
+	{
+		spin_lock(&mylock);
+		WRITE_ONCE(x, 1);
+		WRITE_ONCE(y, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU1(void)
+	{
+		spin_lock(&mylock);
+		smp_mb__after_spinlock();
+		r0 = READ_ONCE(y);
+		WRITE_ONCE(z, 1);
+		spin_unlock(&mylock);
+	}
+
+	void CPU2(void)
+	{
+		WRITE_ONCE(z, 2);
+		smp_mb();
+		r1 = READ_ONCE(x);
+	}
+
+This addition of smp_mb__after_spinlock() strengthens the lock acquisition
+sufficiently to rule out the counter-intuitive outcome.
+
+
+Taking off the training wheels
+==============================
+
+This section looks at more complex examples, including message passing,
+load buffering, release-acquire chains, store buffering.
+Many classes of litmus tests have abbreviated names, which may be found
+here: https://www.cl.cam.ac.uk/~pes20/ppc-supplemental/test6.pdf
+
+
+Message passing (MP)
+--------------------
+
+The MP pattern has one CPU execute a pair of stores to a pair of variables
+and another CPU execute a pair of loads from this same pair of variables,
+but in the opposite order.  The goal is to avoid the counter-intuitive
+outcome in which the first load sees the value written by the second store
+but the second load does not see the value written by the first store.
+In the absence of any ordering, this goal may not be met, as can be seen
+in the MP+poonceonces.litmus litmus test.  This section therefore looks at
+a number of ways of meeting this goal.
+
+
+Release and acquire
+~~~~~~~~~~~~~~~~~~~
+
+Use of smp_store_release() and smp_load_acquire() is one way to force
+the desired MP ordering.  The general approach is shown below:
+
+	/* See MP+pooncerelease+poacquireonce.litmus. */
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		smp_store_release(&y, 1);
+	}
+
+	void CPU1(void)
+	{
+		r0 = smp_load_acquire(&y);
+		r1 = READ_ONCE(x);
+	}
+
+The smp_store_release() macro orders any prior accesses against the
+store, while the smp_load_acquire macro orders the load against any
+subsequent accesses.  Therefore, if the final value of r0 is the value 1,
+the final value of r1 must also be the value 1.
+
+The init_stack_slab() function in lib/stackdepot.c uses release-acquire
+in this way to safely initialize of a slab of the stack.  Working out
+the mutual-exclusion design is left as an exercise for the reader.
+
+
+Assign and dereference
+~~~~~~~~~~~~~~~~~~~~~~
+
+Use of rcu_assign_pointer() and rcu_dereference() is quite similar to the
+use of smp_store_release() and smp_load_acquire(), except that both
+rcu_assign_pointer() and rcu_dereference() operate on RCU-protected
+pointers.  The general approach is shown below:
+
+	/* See MP+onceassign+derefonce.litmus. */
+	int z;
+	int *y = &z;
+	int x;
+
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		rcu_assign_pointer(y, &x);
+	}
+
+	void CPU1(void)
+	{
+		rcu_read_lock();
+		r0 = rcu_dereference(y);
+		r1 = READ_ONCE(*r0);
+		rcu_read_unlock();
+	}
+
+In this example, if the final value of r0 is &x then the final value of
+r1 must be 1.
+
+The rcu_assign_pointer() macro has the same ordering properties as does
+smp_store_release(), but the rcu_dereference() macro orders the load only
+against later accesses that depend on the value loaded.  A dependency
+is present if the value loaded determines the address of a later access
+(address dependency, as shown above), the value written by a later store
+(data dependency), or whether or not a later store is executed in the
+first place (control dependency).  Note that the term "data dependency"
+is sometimes casually used to cover both address and data dependencies.
+
+In lib/math/prime_numbers.c, the expand_to_next_prime() function invokes
+rcu_assign_pointer(), and the next_prime_number() function invokes
+rcu_dereference().  This combination mediates access to a bit vector
+that is expanded as additional primes are needed.
+
+
+Write and read memory barriers
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+It is usually better to use smp_store_release() instead of smp_wmb()
+and to use smp_load_acquire() instead of smp_rmb().  However, the older
+smp_wmb() and smp_rmb() APIs are still heavily used, so it is important
+to understand their use cases.  The general approach is shown below:
+
+	/* See MP+fencewmbonceonce+fencermbonceonce.litmus. */
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		smp_wmb();
+		WRITE_ONCE(y, 1);
+	}
+
+	void CPU1(void)
+	{
+		r0 = READ_ONCE(y);
+		smp_rmb();
+		r1 = READ_ONCE(x);
+	}
+
+The smp_wmb() macro orders prior stores against later stores, and the
+smp_rmb() macro orders prior loads against later loads.  Therefore, if
+the final value of r0 is 1, the final value of r1 must also be 1.
+
+The xlog_state_switch_iclogs() function in fs/xfs/xfs_log.c contains
+the following write-side code fragment:
+
+	log->l_curr_block -= log->l_logBBsize;
+	ASSERT(log->l_curr_block >= 0);
+	smp_wmb();
+	log->l_curr_cycle++;
+
+And the xlog_valid_lsn() function in fs/xfs/xfs_log_priv.h contains
+the corresponding read-side code fragment:
+
+	cur_cycle = READ_ONCE(log->l_curr_cycle);
+	smp_rmb();
+	cur_block = READ_ONCE(log->l_curr_block);
+
+Alternatively, consider the following comment in function
+perf_output_put_handle() in kernel/events/ring_buffer.c:
+
+	 *   kernel				user
+	 *
+	 *   if (LOAD ->data_tail) {		LOAD ->data_head
+	 *			(A)		smp_rmb()	(C)
+	 *	STORE $data			LOAD $data
+	 *	smp_wmb()	(B)		smp_mb()	(D)
+	 *	STORE ->data_head		STORE ->data_tail
+	 *   }
+
+The B/C pairing is an example of the MP pattern using smp_wmb() on the
+write side and smp_rmb() on the read side.
+
+Of course, given that smp_mb() is strictly stronger than either smp_wmb()
+or smp_rmb(), any code fragment that would work with smp_rmb() and
+smp_wmb() would also work with smp_mb() replacing either or both of the
+weaker barriers.
+
+
+Load buffering (LB)
+-------------------
+
+The LB pattern has one CPU load from one variable and then store to a
+second, while another CPU loads from the second variable and then stores
+to the first.  The goal is to avoid the counter-intuitive situation where
+each load reads the value written by the other CPU's store.  In the
+absence of any ordering it is quite possible that this may happen, as
+can be seen in the LB+poonceonces.litmus litmus test.
+
+One way of avoiding the counter-intuitive outcome is through the use of a
+control dependency paired with a full memory barrier:
+
+	/* See LB+fencembonceonce+ctrlonceonce.litmus. */
+	void CPU0(void)
+	{
+		r0 = READ_ONCE(x);
+		if (r0)
+			WRITE_ONCE(y, 1);
+	}
+
+	void CPU1(void)
+	{
+		r1 = READ_ONCE(y);
+		smp_mb();
+		WRITE_ONCE(x, 1);
+	}
+
+This pairing of a control dependency in CPU0() with a full memory
+barrier in CPU1() prevents r0 and r1 from both ending up equal to 1.
+
+The A/D pairing from the ring-buffer use case shown earlier also
+illustrates LB.  Here is a repeat of the comment in
+perf_output_put_handle() in kernel/events/ring_buffer.c, showing a
+control dependency on the kernel side and a full memory barrier on
+the user side:
+
+	 *   kernel				user
+	 *
+	 *   if (LOAD ->data_tail) {		LOAD ->data_head
+	 *			(A)		smp_rmb()	(C)
+	 *	STORE $data			LOAD $data
+	 *	smp_wmb()	(B)		smp_mb()	(D)
+	 *	STORE ->data_head		STORE ->data_tail
+	 *   }
+	 *
+	 * Where A pairs with D, and B pairs with C.
+
+The kernel's control dependency between the load from ->data_tail
+and the store to data combined with the user's full memory barrier
+between the load from data and the store to ->data_tail prevents
+the counter-intuitive outcome where the kernel overwrites the data
+before the user gets done loading it.
+
+
+Release-acquire chains
+----------------------
+
+Release-acquire chains are a low-overhead, flexible, and easy-to-use
+method of maintaining order.  However, they do have some limitations that
+need to be fully understood.  Here is an example that maintains order:
+
+	/* See ISA2+pooncerelease+poacquirerelease+poacquireonce.litmus. */
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		smp_store_release(&y, 1);
+	}
+
+	void CPU1(void)
+	{
+		r0 = smp_load_acquire(y);
+		smp_store_release(&z, 1);
+	}
+
+	void CPU2(void)
+	{
+		r1 = smp_load_acquire(z);
+		r2 = READ_ONCE(x);
+	}
+
+In this case, if r0 and r1 both have final values of 1, then r2 must
+also have a final value of 1.
+
+The ordering in this example is stronger than it needs to be.  For
+example, ordering would still be preserved if CPU1()'s smp_load_acquire()
+invocation was replaced with READ_ONCE().
+
+It is tempting to assume that CPU0()'s store to x is globally ordered
+before CPU1()'s store to z, but this is not the case:
+
+	/* See Z6.0+pooncerelease+poacquirerelease+mbonceonce.litmus. */
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		smp_store_release(&y, 1);
+	}
+
+	void CPU1(void)
+	{
+		r0 = smp_load_acquire(y);
+		smp_store_release(&z, 1);
+	}
+
+	void CPU2(void)
+	{
+		WRITE_ONCE(z, 2);
+		smp_mb();
+		r1 = READ_ONCE(x);
+	}
+
+One might hope that if the final value of r0 is 1 and the final value
+of z is 2, then the final value of r1 must also be 1, but it really is
+possible for r1 to have the final value of 0.  The reason, of course,
+is that in this version, CPU2() is not part of the release-acquire chain.
+This situation is accounted for in the rules of thumb below.
+
+Despite this limitation, release-acquire chains are low-overhead as
+well as simple and powerful, at least as memory-ordering mechanisms go.
+
+
+Store buffering
+---------------
+
+Store buffering can be thought of as upside-down load buffering, so
+that one CPU first stores to one variable and then loads from a second,
+while another CPU stores to the second variable and then loads from the
+first.  Preserving order requires nothing less than full barriers:
+
+	/* See SB+fencembonceonces.litmus. */
+	void CPU0(void)
+	{
+		WRITE_ONCE(x, 1);
+		smp_mb();
+		r0 = READ_ONCE(y);
+	}
+
+	void CPU1(void)
+	{
+		WRITE_ONCE(y, 1);
+		smp_mb();
+		r1 = READ_ONCE(x);
+	}
+
+Omitting either smp_mb() will allow both r0 and r1 to have final
+values of 0, but providing both full barriers as shown above prevents
+this counter-intuitive outcome.
+
+This pattern most famously appears as part of Dekker's locking
+algorithm, but it has a much more practical use within the Linux kernel
+of ordering wakeups.  The following comment taken from waitqueue_active()
+in include/linux/wait.h shows the canonical pattern:
+
+ *      CPU0 - waker                    CPU1 - waiter
+ *
+ *                                      for (;;) {
+ *      @cond = true;                     prepare_to_wait(&wq_head, &wait, state);
+ *      smp_mb();                         // smp_mb() from set_current_state()
+ *      if (waitqueue_active(wq_head))         if (@cond)
+ *        wake_up(wq_head);                      break;
+ *                                        schedule();
+ *                                      }
+ *                                      finish_wait(&wq_head, &wait);
+
+On CPU0, the store is to @cond and the load is in waitqueue_active().
+On CPU1, prepare_to_wait() contains both a store to wq_head and a call
+to set_current_state(), which contains an smp_mb() barrier; the load is
+"if (@cond)".  The full barriers prevent the undesirable outcome where
+CPU1 puts the waiting task to sleep and CPU0 fails to wake it up.
+
+Note that use of locking can greatly simplify this pattern.
+
+
+Rules of thumb
+==============
+
+There might seem to be no pattern governing what ordering primitives are
+needed in which situations, but this is not the case.  There is a pattern
+based on the relation between the accesses linking successive CPUs in a
+given litmus test.  There are three types of linkage:
+
+1.	Write-to-read, where the next CPU reads the value that the
+	previous CPU wrote.  The LB litmus-test patterns contain only
+	this type of relation.	In formal memory-modeling texts, this
+	relation is called "reads-from" and is usually abbreviated "rf".
+
+2.	Read-to-write, where the next CPU overwrites the value that the
+	previous CPU read.  The SB litmus test contains only this type
+	of relation.  In formal memory-modeling texts, this relation is
+	often called "from-reads" and is sometimes abbreviated "fr".
+
+3.	Write-to-write, where the next CPU overwrites the value written
+	by the previous CPU.  The Z6.0 litmus test pattern contains a
+	write-to-write relation between the last access of CPU1() and
+	the first access of CPU2().  In formal memory-modeling texts,
+	this relation is often called "coherence order" and is sometimes
+	abbreviated "co".  In the C++ standard, it is instead called
+	"modification order" and often abbreviated "mo".
+
+The strength of memory ordering required for a given litmus test to
+avoid a counter-intuitive outcome depends on the types of relations
+linking the memory accesses for the outcome in question:
+
+o	If all links are write-to-read links, then the weakest
+	possible ordering within each CPU suffices.  For example, in
+	the LB litmus test, a control dependency was enough to do the
+	job.
+
+o	If all but one of the links are write-to-read links, then a
+	release-acquire chain suffices.  Both the MP and the ISA2
+	litmus tests illustrate this case.
+
+o	If more than one of the links are something other than
+	write-to-read links, then a full memory barrier is required
+	between each successive pair of non-write-to-read links.  This
+	case is illustrated by the Z6.0 litmus tests, both in the
+	locking and in the release-acquire sections.
+
+However, if you find yourself having to stretch these rules of thumb
+to fit your situation, you should consider creating a litmus test and
+running it on the model.
diff --git a/tools/memory-model/Documentation/references.txt b/tools/memory-model/Documentation/references.txt
new file mode 100644
index 0000000000..c5fdfd19df
--- /dev/null
+++ b/tools/memory-model/Documentation/references.txt
@@ -0,0 +1,131 @@
+This document provides background reading for memory models and related
+tools.  These documents are aimed at kernel hackers who are interested
+in memory models.
+
+
+Hardware manuals and models
+===========================
+
+o	SPARC International Inc. (Ed.). 1994. "The SPARC Architecture
+	Reference Manual Version 9". SPARC International Inc.
+
+o	Compaq Computer Corporation (Ed.). 2002. "Alpha Architecture
+	Reference Manual".  Compaq Computer Corporation.
+
+o	Intel Corporation (Ed.). 2002. "A Formal Specification of Intel
+	Itanium Processor Family Memory Ordering". Intel Corporation.
+
+o	Intel Corporation (Ed.). 2002. "Intel 64 and IA-32 Architectures
+	Software Developer’s Manual". Intel Corporation.
+
+o	Peter Sewell, Susmit Sarkar, Scott Owens, Francesco Zappa Nardelli,
+	and Magnus O. Myreen. 2010. "x86-TSO: A Rigorous and Usable
+	Programmer's Model for x86 Multiprocessors". Commun. ACM 53, 7
+	(July, 2010), 89-97. http://doi.acm.org/10.1145/1785414.1785443
+
+o	IBM Corporation (Ed.). 2009. "Power ISA Version 2.06". IBM
+	Corporation.
+
+o	ARM Ltd. (Ed.). 2009. "ARM Barrier Litmus Tests and Cookbook".
+	ARM Ltd.
+
+o	Susmit Sarkar, Peter Sewell, Jade Alglave, Luc Maranget, and
+	Derek Williams.  2011. "Understanding POWER Multiprocessors". In
+	Proceedings of the 32Nd ACM SIGPLAN Conference on Programming
+	Language Design and Implementation (PLDI ’11). ACM, New York,
+	NY, USA, 175–186.
+
+o	Susmit Sarkar, Kayvan Memarian, Scott Owens, Mark Batty,
+	Peter Sewell, Luc Maranget, Jade Alglave, and Derek Williams.
+	2012. "Synchronising C/C++ and POWER". In Proceedings of the 33rd
+	ACM SIGPLAN Conference on Programming Language Design and
+	Implementation (PLDI '12). ACM, New York, NY, USA, 311-322.
+
+o	ARM Ltd. (Ed.). 2014. "ARM Architecture Reference Manual (ARMv8,
+	for ARMv8-A architecture profile)". ARM Ltd.
+
+o	Imagination Technologies, LTD. 2015. "MIPS(R) Architecture
+	For Programmers, Volume II-A: The MIPS64(R) Instruction,
+	Set Reference Manual". Imagination Technologies,
+	LTD. https://imgtec.com/?do-download=4302.
+
+o	Shaked Flur, Kathryn E. Gray, Christopher Pulte, Susmit
+	Sarkar, Ali Sezgin, Luc Maranget, Will Deacon, and Peter
+	Sewell. 2016. "Modelling the ARMv8 Architecture, Operationally:
+	Concurrency and ISA". In Proceedings of the 43rd Annual ACM
+	SIGPLAN-SIGACT Symposium on Principles of Programming Languages
+	(POPL ’16). ACM, New York, NY, USA, 608–621.
+
+o	Shaked Flur, Susmit Sarkar, Christopher Pulte, Kyndylan Nienhuis,
+	Luc Maranget, Kathryn E. Gray, Ali Sezgin, Mark Batty, and Peter
+	Sewell. 2017. "Mixed-size Concurrency: ARM, POWER, C/C++11,
+	and SC". In Proceedings of the 44th ACM SIGPLAN Symposium on
+	Principles of Programming Languages (POPL 2017). ACM, New York,
+	NY, USA, 429–442.
+
+o	Christopher Pulte, Shaked Flur, Will Deacon, Jon French,
+	Susmit Sarkar, and Peter Sewell. 2018. "Simplifying ARM concurrency:
+	multicopy-atomic axiomatic and operational models for ARMv8". In
+	Proceedings of the ACM on Programming Languages, Volume 2, Issue
+	POPL, Article No. 19. ACM, New York, NY, USA.
+
+
+Linux-kernel memory model
+=========================
+
+o	Jade Alglave, Will Deacon, Boqun Feng, David Howells, Daniel
+	Lustig, Luc Maranget, Paul E. McKenney, Andrea Parri, Nicholas
+	Piggin, Alan Stern, Akira Yokosawa, and Peter Zijlstra.
+	2019. "Calibrating your fear of big bad optimizing compilers"
+	Linux Weekly News.  https://lwn.net/Articles/799218/
+
+o	Jade Alglave, Will Deacon, Boqun Feng, David Howells, Daniel
+	Lustig, Luc Maranget, Paul E. McKenney, Andrea Parri, Nicholas
+	Piggin, Alan Stern, Akira Yokosawa, and Peter Zijlstra.
+	2019. "Who's afraid of a big bad optimizing compiler?"
+	Linux Weekly News.  https://lwn.net/Articles/793253/
+
+o	Jade Alglave, Luc Maranget, Paul E. McKenney, Andrea Parri, and
+	Alan Stern.  2018. "Frightening small children and disconcerting
+	grown-ups: Concurrency in the Linux kernel". In Proceedings of
+	the 23rd International Conference on Architectural Support for
+	Programming Languages and Operating Systems (ASPLOS 2018). ACM,
+	New York, NY, USA, 405-418.  Webpage: http://diy.inria.fr/linux/.
+
+o	Jade Alglave, Luc Maranget, Paul E. McKenney, Andrea Parri, and
+	Alan Stern.  2017.  "A formal kernel memory-ordering model (part 1)"
+	Linux Weekly News.  https://lwn.net/Articles/718628/
+
+o	Jade Alglave, Luc Maranget, Paul E. McKenney, Andrea Parri, and
+	Alan Stern.  2017.  "A formal kernel memory-ordering model (part 2)"
+	Linux Weekly News.  https://lwn.net/Articles/720550/
+
+o	Jade Alglave, Luc Maranget, Paul E. McKenney, Andrea Parri, and
+	Alan Stern.  2017-2019.  "A Formal Model of Linux-Kernel Memory
+	Ordering" (backup material for the LWN articles)
+	https://mirrors.edge.kernel.org/pub/linux/kernel/people/paulmck/LWNLinuxMM/
+
+
+Memory-model tooling
+====================
+
+o	Daniel Jackson. 2002. "Alloy: A Lightweight Object Modelling
+	Notation". ACM Trans. Softw. Eng. Methodol. 11, 2 (April 2002),
+	256–290. http://doi.acm.org/10.1145/505145.505149
+
+o	Jade Alglave, Luc Maranget, and Michael Tautschnig. 2014. "Herding
+	Cats: Modelling, Simulation, Testing, and Data Mining for Weak
+	Memory". ACM Trans. Program. Lang. Syst. 36, 2, Article 7 (July
+	2014), 7:1–7:74 pages.
+
+o	Jade Alglave, Patrick Cousot, and Luc Maranget. 2016. "Syntax and
+	semantics of the weak consistency model specification language
+	cat". CoRR abs/1608.07531 (2016). https://arxiv.org/abs/1608.07531
+
+
+Memory-model comparisons
+========================
+
+o	Paul E. McKenney, Ulrich Weigand, Andrea Parri, and Boqun
+	Feng. 2018. "Linux-Kernel Memory Model". (27 September 2018).
+	http://www.open-std.org/jtc1/sc22/wg21/docs/papers/2018/p0124r6.html.
diff --git a/tools/memory-model/Documentation/simple.txt b/tools/memory-model/Documentation/simple.txt
new file mode 100644
index 0000000000..4c789ec833
--- /dev/null
+++ b/tools/memory-model/Documentation/simple.txt
@@ -0,0 +1,270 @@
+This document provides options for those wishing to keep their
+memory-ordering lives simple, as is necessary for those whose domain
+is complex.  After all, there are bugs other than memory-ordering bugs,
+and the time spent gaining memory-ordering knowledge is not available
+for gaining domain knowledge.  Furthermore Linux-kernel memory model
+(LKMM) is quite complex, with subtle differences in code often having
+dramatic effects on correctness.
+
+The options near the beginning of this list are quite simple.  The idea
+is not that kernel hackers don't already know about them, but rather
+that they might need the occasional reminder.
+
+Please note that this is a generic guide, and that specific subsystems
+will often have special requirements or idioms.  For example, developers
+of MMIO-based device drivers will often need to use mb(), rmb(), and
+wmb(), and therefore might find smp_mb(), smp_rmb(), and smp_wmb()
+to be more natural than smp_load_acquire() and smp_store_release().
+On the other hand, those coming in from other environments will likely
+be more familiar with these last two.
+
+
+Single-threaded code
+====================
+
+In single-threaded code, there is no reordering, at least assuming
+that your toolchain and hardware are working correctly.  In addition,
+it is generally a mistake to assume your code will only run in a single
+threaded context as the kernel can enter the same code path on multiple
+CPUs at the same time.  One important exception is a function that makes
+no external data references.
+
+In the general case, you will need to take explicit steps to ensure that
+your code really is executed within a single thread that does not access
+shared variables.  A simple way to achieve this is to define a global lock
+that you acquire at the beginning of your code and release at the end,
+taking care to ensure that all references to your code's shared data are
+also carried out under that same lock.  Because only one thread can hold
+this lock at a given time, your code will be executed single-threaded.
+This approach is called "code locking".
+
+Code locking can severely limit both performance and scalability, so it
+should be used with caution, and only on code paths that execute rarely.
+After all, a huge amount of effort was required to remove the Linux
+kernel's old "Big Kernel Lock", so let's please be very careful about
+adding new "little kernel locks".
+
+One of the advantages of locking is that, in happy contrast with the
+year 1981, almost all kernel developers are very familiar with locking.
+The Linux kernel's lockdep (CONFIG_PROVE_LOCKING=y) is very helpful with
+the formerly feared deadlock scenarios.
+
+Please use the standard locking primitives provided by the kernel rather
+than rolling your own.  For one thing, the standard primitives interact
+properly with lockdep.  For another thing, these primitives have been
+tuned to deal better with high contention.  And for one final thing, it is
+surprisingly hard to correctly code production-quality lock acquisition
+and release functions.  After all, even simple non-production-quality
+locking functions must carefully prevent both the CPU and the compiler
+from moving code in either direction across the locking function.
+
+Despite the scalability limitations of single-threaded code, RCU
+takes this approach for much of its grace-period processing and also
+for early-boot operation.  The reason RCU is able to scale despite
+single-threaded grace-period processing is use of batching, where all
+updates that accumulated during one grace period are handled by the
+next one.  In other words, slowing down grace-period processing makes
+it more efficient.  Nor is RCU unique:  Similar batching optimizations
+are used in many I/O operations.
+
+
+Packaged code
+=============
+
+Even if performance and scalability concerns prevent your code from
+being completely single-threaded, it is often possible to use library
+functions that handle the concurrency nearly or entirely on their own.
+This approach delegates any LKMM worries to the library maintainer.
+
+In the kernel, what is the "library"?  Quite a bit.  It includes the
+contents of the lib/ directory, much of the include/linux/ directory along
+with a lot of other heavily used APIs.  But heavily used examples include
+the list macros (for example, include/linux/{,rcu}list.h), workqueues,
+smp_call_function(), and the various hash tables and search trees.
+
+
+Data locking
+============
+
+With code locking, we use single-threaded code execution to guarantee
+serialized access to the data that the code is accessing.  However,
+we can also achieve this by instead associating the lock with specific
+instances of the data structures.  This creates a "critical section"
+in the code execution that will execute as though it is single threaded.
+By placing all the accesses and modifications to a shared data structure
+inside a critical section, we ensure that the execution context that
+holds the lock has exclusive access to the shared data.
+
+The poster boy for this approach is the hash table, where placing a lock
+in each hash bucket allows operations on different buckets to proceed
+concurrently.  This works because the buckets do not overlap with each
+other, so that an operation on one bucket does not interfere with any
+other bucket.
+
+As the number of buckets increases, data locking scales naturally.
+In particular, if the amount of data increases with the number of CPUs,
+increasing the number of buckets as the number of CPUs increase results
+in a naturally scalable data structure.
+
+
+Per-CPU processing
+==================
+
+Partitioning processing and data over CPUs allows each CPU to take
+a single-threaded approach while providing excellent performance and
+scalability.  Of course, there is no free lunch:  The dark side of this
+excellence is substantially increased memory footprint.
+
+In addition, it is sometimes necessary to occasionally update some global
+view of this processing and data, in which case something like locking
+must be used to protect this global view.  This is the approach taken
+by the percpu_counter infrastructure. In many cases, there are already
+generic/library variants of commonly used per-cpu constructs available.
+Please use them rather than rolling your own.
+
+RCU uses DEFINE_PER_CPU*() declaration to create a number of per-CPU
+data sets.  For example, each CPU does private quiescent-state processing
+within its instance of the per-CPU rcu_data structure, and then uses data
+locking to report quiescent states up the grace-period combining tree.
+
+
+Packaged primitives: Sequence locking
+=====================================
+
+Lockless programming is considered by many to be more difficult than
+lock-based programming, but there are a few lockless design patterns that
+have been built out into an API.  One of these APIs is sequence locking.
+Although this APIs can be used in extremely complex ways, there are simple
+and effective ways of using it that avoid the need to pay attention to
+memory ordering.
+
+The basic keep-things-simple rule for sequence locking is "do not write
+in read-side code".  Yes, you can do writes from within sequence-locking
+readers, but it won't be so simple.  For example, such writes will be
+lockless and should be idempotent.
+
+For more sophisticated use cases, LKMM can guide you, including use
+cases involving combining sequence locking with other synchronization
+primitives.  (LKMM does not yet know about sequence locking, so it is
+currently necessary to open-code it in your litmus tests.)
+
+Additional information may be found in include/linux/seqlock.h.
+
+Packaged primitives: RCU
+========================
+
+Another lockless design pattern that has been baked into an API
+is RCU.  The Linux kernel makes sophisticated use of RCU, but the
+keep-things-simple rules for RCU are "do not write in read-side code"
+and "do not update anything that is visible to and accessed by readers",
+and "protect updates with locking".
+
+These rules are illustrated by the functions foo_update_a() and
+foo_get_a() shown in Documentation/RCU/whatisRCU.rst.  Additional
+RCU usage patterns maybe found in Documentation/RCU and in the
+source code.
+
+
+Packaged primitives: Atomic operations
+======================================
+
+Back in the day, the Linux kernel had three types of atomic operations:
+
+1.	Initialization and read-out, such as atomic_set() and atomic_read().
+
+2.	Operations that did not return a value and provided no ordering,
+	such as atomic_inc() and atomic_dec().
+
+3.	Operations that returned a value and provided full ordering, such as
+	atomic_add_return() and atomic_dec_and_test().  Note that some
+	value-returning operations provide full ordering only conditionally.
+	For example, cmpxchg() provides ordering only upon success.
+
+More recent kernels have operations that return a value but do not
+provide full ordering.  These are flagged with either a _relaxed()
+suffix (providing no ordering), or an _acquire() or _release() suffix
+(providing limited ordering).
+
+Additional information may be found in these files:
+
+Documentation/atomic_t.txt
+Documentation/atomic_bitops.txt
+Documentation/core-api/refcount-vs-atomic.rst
+
+Reading code using these primitives is often also quite helpful.
+
+
+Lockless, fully ordered
+=======================
+
+When using locking, there often comes a time when it is necessary
+to access some variable or another without holding the data lock
+that serializes access to that variable.
+
+If you want to keep things simple, use the initialization and read-out
+operations from the previous section only when there are no racing
+accesses.  Otherwise, use only fully ordered operations when accessing
+or modifying the variable.  This approach guarantees that code prior
+to a given access to that variable will be seen by all CPUs has having
+happened before any code following any later access to that same variable.
+
+Please note that per-CPU functions are not atomic operations and
+hence they do not provide any ordering guarantees at all.
+
+If the lockless accesses are frequently executed reads that are used
+only for heuristics, or if they are frequently executed writes that
+are used only for statistics, please see the next section.
+
+
+Lockless statistics and heuristics
+==================================
+
+Unordered primitives such as atomic_read(), atomic_set(), READ_ONCE(), and
+WRITE_ONCE() can safely be used in some cases.  These primitives provide
+no ordering, but they do prevent the compiler from carrying out a number
+of destructive optimizations (for which please see the next section).
+One example use for these primitives is statistics, such as per-CPU
+counters exemplified by the rt_cache_stat structure's routing-cache
+statistics counters.  Another example use case is heuristics, such as
+the jiffies_till_first_fqs and jiffies_till_next_fqs kernel parameters
+controlling how often RCU scans for idle CPUs.
+
+But be careful.  "Unordered" really does mean "unordered".  It is all
+too easy to assume ordering, and this assumption must be avoided when
+using these primitives.
+
+
+Don't let the compiler trip you up
+==================================
+
+It can be quite tempting to use plain C-language accesses for lockless
+loads from and stores to shared variables.  Although this is both
+possible and quite common in the Linux kernel, it does require a
+surprising amount of analysis, care, and knowledge about the compiler.
+Yes, some decades ago it was not unfair to consider a C compiler to be
+an assembler with added syntax and better portability, but the advent of
+sophisticated optimizing compilers mean that those days are long gone.
+Today's optimizing compilers can profoundly rewrite your code during the
+translation process, and have long been ready, willing, and able to do so.
+
+Therefore, if you really need to use C-language assignments instead of
+READ_ONCE(), WRITE_ONCE(), and so on, you will need to have a very good
+understanding of both the C standard and your compiler.  Here are some
+introductory references and some tooling to start you on this noble quest:
+
+Who's afraid of a big bad optimizing compiler?
+	https://lwn.net/Articles/793253/
+Calibrating your fear of big bad optimizing compilers
+	https://lwn.net/Articles/799218/
+Concurrency bugs should fear the big bad data-race detector (part 1)
+	https://lwn.net/Articles/816850/
+Concurrency bugs should fear the big bad data-race detector (part 2)
+	https://lwn.net/Articles/816854/
+
+
+More complex use cases
+======================
+
+If the alternatives above do not do what you need, please look at the
+recipes-pairs.txt file to peel off the next layer of the memory-ordering
+onion.
diff --git a/tools/memory-model/README b/tools/memory-model/README
new file mode 100644
index 0000000000..dab3890420
--- /dev/null
+++ b/tools/memory-model/README
@@ -0,0 +1,223 @@
+		=====================================
+		LINUX KERNEL MEMORY CONSISTENCY MODEL
+		=====================================
+
+============
+INTRODUCTION
+============
+
+This directory contains the memory consistency model (memory model, for
+short) of the Linux kernel, written in the "cat" language and executable
+by the externally provided "herd7" simulator, which exhaustively explores
+the state space of small litmus tests.
+
+In addition, the "klitmus7" tool (also externally provided) may be used
+to convert a litmus test to a Linux kernel module, which in turn allows
+that litmus test to be exercised within the Linux kernel.
+
+
+============
+REQUIREMENTS
+============
+
+Version 7.52 or higher of the "herd7" and "klitmus7" tools must be
+downloaded separately:
+
+  https://github.com/herd/herdtools7
+
+See "herdtools7/INSTALL.md" for installation instructions.
+
+Note that although these tools usually provide backwards compatibility,
+this is not absolutely guaranteed.
+
+For example, a future version of herd7 might not work with the model
+in this release.  A compatible model will likely be made available in
+a later release of Linux kernel.
+
+If you absolutely need to run the model in this particular release,
+please try using the exact version called out above.
+
+klitmus7 is independent of the model provided here.  It has its own
+dependency on a target kernel release where converted code is built
+and executed.  Any change in kernel APIs essential to klitmus7 will
+necessitate an upgrade of klitmus7.
+
+If you find any compatibility issues in klitmus7, please inform the
+memory model maintainers.
+
+klitmus7 Compatibility Table
+----------------------------
+
+	============  ==========
+	target Linux  herdtools7
+	------------  ----------
+	     -- 4.14  7.48 --
+	4.15 -- 4.19  7.49 --
+	4.20 -- 5.5   7.54 --
+	5.6  -- 5.16  7.56 --
+	5.17 --       7.56.1 --
+	============  ==========
+
+
+==================
+BASIC USAGE: HERD7
+==================
+
+The memory model is used, in conjunction with "herd7", to exhaustively
+explore the state space of small litmus tests.  Documentation describing
+the format, features, capabilities and limitations of these litmus
+tests is available in tools/memory-model/Documentation/litmus-tests.txt.
+
+Example litmus tests may be found in the Linux-kernel source tree:
+
+	tools/memory-model/litmus-tests/
+	Documentation/litmus-tests/
+
+Several thousand more example litmus tests are available here:
+
+	https://github.com/paulmckrcu/litmus
+	https://git.kernel.org/pub/scm/linux/kernel/git/paulmck/perfbook.git/tree/CodeSamples/formal/herd
+	https://git.kernel.org/pub/scm/linux/kernel/git/paulmck/perfbook.git/tree/CodeSamples/formal/litmus
+
+Documentation describing litmus tests and now to use them may be found
+here:
+
+	tools/memory-model/Documentation/litmus-tests.txt
+
+The remainder of this section uses the SB+fencembonceonces.litmus test
+located in the tools/memory-model directory.
+
+To run SB+fencembonceonces.litmus against the memory model:
+
+  $ cd $LINUX_SOURCE_TREE/tools/memory-model
+  $ herd7 -conf linux-kernel.cfg litmus-tests/SB+fencembonceonces.litmus
+
+Here is the corresponding output:
+
+  Test SB+fencembonceonces Allowed
+  States 3
+  0:r0=0; 1:r0=1;
+  0:r0=1; 1:r0=0;
+  0:r0=1; 1:r0=1;
+  No
+  Witnesses
+  Positive: 0 Negative: 3
+  Condition exists (0:r0=0 /\ 1:r0=0)
+  Observation SB+fencembonceonces Never 0 3
+  Time SB+fencembonceonces 0.01
+  Hash=d66d99523e2cac6b06e66f4c995ebb48
+
+The "Positive: 0 Negative: 3" and the "Never 0 3" each indicate that
+this litmus test's "exists" clause can not be satisfied.
+
+See "herd7 -help" or "herdtools7/doc/" for more information on running the
+tool itself, but please be aware that this documentation is intended for
+people who work on the memory model itself, that is, people making changes
+to the tools/memory-model/linux-kernel.* files.  It is not intended for
+people focusing on writing, understanding, and running LKMM litmus tests.
+
+
+=====================
+BASIC USAGE: KLITMUS7
+=====================
+
+The "klitmus7" tool converts a litmus test into a Linux kernel module,
+which may then be loaded and run.
+
+For example, to run SB+fencembonceonces.litmus against hardware:
+
+  $ mkdir mymodules
+  $ klitmus7 -o mymodules litmus-tests/SB+fencembonceonces.litmus
+  $ cd mymodules ; make
+  $ sudo sh run.sh
+
+The corresponding output includes:
+
+  Test SB+fencembonceonces Allowed
+  Histogram (3 states)
+  644580  :>0:r0=1; 1:r0=0;
+  644328  :>0:r0=0; 1:r0=1;
+  711092  :>0:r0=1; 1:r0=1;
+  No
+  Witnesses
+  Positive: 0, Negative: 2000000
+  Condition exists (0:r0=0 /\ 1:r0=0) is NOT validated
+  Hash=d66d99523e2cac6b06e66f4c995ebb48
+  Observation SB+fencembonceonces Never 0 2000000
+  Time SB+fencembonceonces 0.16
+
+The "Positive: 0 Negative: 2000000" and the "Never 0 2000000" indicate
+that during two million trials, the state specified in this litmus
+test's "exists" clause was not reached.
+
+And, as with "herd7", please see "klitmus7 -help" or "herdtools7/doc/"
+for more information.  And again, please be aware that this documentation
+is intended for people who work on the memory model itself, that is,
+people making changes to the tools/memory-model/linux-kernel.* files.
+It is not intended for people focusing on writing, understanding, and
+running LKMM litmus tests.
+
+
+====================
+DESCRIPTION OF FILES
+====================
+
+Documentation/README
+	Guide to the other documents in the Documentation/ directory.
+
+linux-kernel.bell
+	Categorizes the relevant instructions, including memory
+	references, memory barriers, atomic read-modify-write operations,
+	lock acquisition/release, and RCU operations.
+
+	More formally, this file (1) lists the subtypes of the various
+	event types used by the memory model and (2) performs RCU
+	read-side critical section nesting analysis.
+
+linux-kernel.cat
+	Specifies what reorderings are forbidden by memory references,
+	memory barriers, atomic read-modify-write operations, and RCU.
+
+	More formally, this file specifies what executions are forbidden
+	by the memory model.  Allowed executions are those which
+	satisfy the model's "coherence", "atomic", "happens-before",
+	"propagation", and "rcu" axioms, which are defined in the file.
+
+linux-kernel.cfg
+	Convenience file that gathers the common-case herd7 command-line
+	arguments.
+
+linux-kernel.def
+	Maps from C-like syntax to herd7's internal litmus-test
+	instruction-set architecture.
+
+litmus-tests
+	Directory containing a few representative litmus tests, which
+	are listed in litmus-tests/README.  A great deal more litmus
+	tests are available at https://github.com/paulmckrcu/litmus.
+
+	By "representative", it means the one in the litmus-tests
+	directory is:
+
+		1) simple, the number of threads should be relatively
+		   small and each thread function should be relatively
+		   simple.
+		2) orthogonal, there should be no two litmus tests
+		   describing the same aspect of the memory model.
+		3) textbook, developers can easily copy-paste-modify
+		   the litmus tests to use the patterns on their own
+		   code.
+
+lock.cat
+	Provides a front-end analysis of lock acquisition and release,
+	for example, associating a lock acquisition with the preceding
+	and following releases and checking for self-deadlock.
+
+	More formally, this file defines a performance-enhanced scheme
+	for generation of the possible reads-from and coherence order
+	relations on the locking primitives.
+
+README
+	This file.
+
+scripts	Various scripts, see scripts/README.
diff --git a/tools/memory-model/linux-kernel.bell b/tools/memory-model/linux-kernel.bell
new file mode 100644
index 0000000000..ce06870093
--- /dev/null
+++ b/tools/memory-model/linux-kernel.bell
@@ -0,0 +1,84 @@
+// SPDX-License-Identifier: GPL-2.0+
+(*
+ * Copyright (C) 2015 Jade Alglave <j.alglave@ucl.ac.uk>,
+ * Copyright (C) 2016 Luc Maranget <luc.maranget@inria.fr> for Inria
+ * Copyright (C) 2017 Alan Stern <stern@rowland.harvard.edu>,
+ *                    Andrea Parri <parri.andrea@gmail.com>
+ *
+ * An earlier version of this file appeared in the companion webpage for
+ * "Frightening small children and disconcerting grown-ups: Concurrency
+ * in the Linux kernel" by Alglave, Maranget, McKenney, Parri, and Stern,
+ * which appeared in ASPLOS 2018.
+ *)
+
+"Linux-kernel memory consistency model"
+
+enum Accesses = 'once (*READ_ONCE,WRITE_ONCE*) ||
+		'release (*smp_store_release*) ||
+		'acquire (*smp_load_acquire*) ||
+		'noreturn (* R of non-return RMW *)
+instructions R[{'once,'acquire,'noreturn}]
+instructions W[{'once,'release}]
+instructions RMW[{'once,'acquire,'release}]
+
+enum Barriers = 'wmb (*smp_wmb*) ||
+		'rmb (*smp_rmb*) ||
+		'mb (*smp_mb*) ||
+		'barrier (*barrier*) ||
+		'rcu-lock (*rcu_read_lock*)  ||
+		'rcu-unlock (*rcu_read_unlock*) ||
+		'sync-rcu (*synchronize_rcu*) ||
+		'before-atomic (*smp_mb__before_atomic*) ||
+		'after-atomic (*smp_mb__after_atomic*) ||
+		'after-spinlock (*smp_mb__after_spinlock*) ||
+		'after-unlock-lock (*smp_mb__after_unlock_lock*) ||
+		'after-srcu-read-unlock (*smp_mb__after_srcu_read_unlock*)
+instructions F[Barriers]
+
+(* SRCU *)
+enum SRCU = 'srcu-lock || 'srcu-unlock || 'sync-srcu
+instructions SRCU[SRCU]
+(* All srcu events *)
+let Srcu = Srcu-lock | Srcu-unlock | Sync-srcu
+
+(* Compute matching pairs of nested Rcu-lock and Rcu-unlock *)
+let rcu-rscs = let rec
+	    unmatched-locks = Rcu-lock \ domain(matched)
+	and unmatched-unlocks = Rcu-unlock \ range(matched)
+	and unmatched = unmatched-locks | unmatched-unlocks
+	and unmatched-po = [unmatched] ; po ; [unmatched]
+	and unmatched-locks-to-unlocks =
+		[unmatched-locks] ; po ; [unmatched-unlocks]
+	and matched = matched | (unmatched-locks-to-unlocks \
+		(unmatched-po ; unmatched-po))
+	in matched
+
+(* Validate nesting *)
+flag ~empty Rcu-lock \ domain(rcu-rscs) as unmatched-rcu-lock
+flag ~empty Rcu-unlock \ range(rcu-rscs) as unmatched-rcu-unlock
+
+(* Compute matching pairs of nested Srcu-lock and Srcu-unlock *)
+let carry-srcu-data = (data ; [~ Srcu-unlock] ; rf)*
+let srcu-rscs = ([Srcu-lock] ; carry-srcu-data ; data ; [Srcu-unlock]) & loc
+
+(* Validate nesting *)
+flag ~empty Srcu-lock \ domain(srcu-rscs) as unmatched-srcu-lock
+flag ~empty Srcu-unlock \ range(srcu-rscs) as unmatched-srcu-unlock
+flag ~empty (srcu-rscs^-1 ; srcu-rscs) \ id as multiple-srcu-matches
+
+(* Check for use of synchronize_srcu() inside an RCU critical section *)
+flag ~empty rcu-rscs & (po ; [Sync-srcu] ; po) as invalid-sleep
+
+(* Validate SRCU dynamic match *)
+flag ~empty different-values(srcu-rscs) as srcu-bad-value-match
+
+(* Compute marked and plain memory accesses *)
+let Marked = (~M) | IW | Once | Release | Acquire | domain(rmw) | range(rmw) |
+		LKR | LKW | UL | LF | RL | RU | Srcu-lock | Srcu-unlock
+let Plain = M \ Marked
+
+(* Redefine dependencies to include those carried through plain accesses *)
+let carry-dep = (data ; [~ Srcu-unlock] ; rfi)*
+let addr = carry-dep ; addr
+let ctrl = carry-dep ; ctrl
+let data = carry-dep ; data
diff --git a/tools/memory-model/linux-kernel.cat b/tools/memory-model/linux-kernel.cat
new file mode 100644
index 0000000000..adf3c4f412
--- /dev/null
+++ b/tools/memory-model/linux-kernel.cat
@@ -0,0 +1,216 @@
+// SPDX-License-Identifier: GPL-2.0+
+(*
+ * Copyright (C) 2015 Jade Alglave <j.alglave@ucl.ac.uk>,
+ * Copyright (C) 2016 Luc Maranget <luc.maranget@inria.fr> for Inria
+ * Copyright (C) 2017 Alan Stern <stern@rowland.harvard.edu>,
+ *                    Andrea Parri <parri.andrea@gmail.com>
+ *
+ * An earlier version of this file appeared in the companion webpage for
+ * "Frightening small children and disconcerting grown-ups: Concurrency
+ * in the Linux kernel" by Alglave, Maranget, McKenney, Parri, and Stern,
+ * which appeared in ASPLOS 2018.
+ *)
+
+"Linux-kernel memory consistency model"
+
+(*
+ * File "lock.cat" handles locks and is experimental.
+ * It can be replaced by include "cos.cat" for tests that do not use locks.
+ *)
+
+include "lock.cat"
+
+(*******************)
+(* Basic relations *)
+(*******************)
+
+(* Release Acquire *)
+let acq-po = [Acquire] ; po ; [M]
+let po-rel = [M] ; po ; [Release]
+let po-unlock-lock-po = po ; [UL] ; (po|rf) ; [LKR] ; po
+
+(* Fences *)
+let R4rmb = R \ Noreturn	(* Reads for which rmb works *)
+let rmb = [R4rmb] ; fencerel(Rmb) ; [R4rmb]
+let wmb = [W] ; fencerel(Wmb) ; [W]
+let mb = ([M] ; fencerel(Mb) ; [M]) |
+	([M] ; fencerel(Before-atomic) ; [RMW] ; po? ; [M]) |
+	([M] ; po? ; [RMW] ; fencerel(After-atomic) ; [M]) |
+	([M] ; po? ; [LKW] ; fencerel(After-spinlock) ; [M]) |
+(*
+ * Note: The po-unlock-lock-po relation only passes the lock to the direct
+ * successor, perhaps giving the impression that the ordering of the
+ * smp_mb__after_unlock_lock() fence only affects a single lock handover.
+ * However, in a longer sequence of lock handovers, the implicit
+ * A-cumulative release fences of lock-release ensure that any stores that
+ * propagate to one of the involved CPUs before it hands over the lock to
+ * the next CPU will also propagate to the final CPU handing over the lock
+ * to the CPU that executes the fence.  Therefore, all those stores are
+ * also affected by the fence.
+ *)
+	([M] ; po-unlock-lock-po ;
+		[After-unlock-lock] ; po ; [M]) |
+	([M] ; po? ; [Srcu-unlock] ; fencerel(After-srcu-read-unlock) ; [M])
+let gp = po ; [Sync-rcu | Sync-srcu] ; po?
+let strong-fence = mb | gp
+
+let nonrw-fence = strong-fence | po-rel | acq-po
+let fence = nonrw-fence | wmb | rmb
+let barrier = fencerel(Barrier | Rmb | Wmb | Mb | Sync-rcu | Sync-srcu |
+		Before-atomic | After-atomic | Acquire | Release |
+		Rcu-lock | Rcu-unlock | Srcu-lock | Srcu-unlock) |
+	(po ; [Release]) | ([Acquire] ; po)
+
+(**********************************)
+(* Fundamental coherence ordering *)
+(**********************************)
+
+(* Sequential Consistency Per Variable *)
+let com = rf | co | fr
+acyclic po-loc | com as coherence
+
+(* Atomic Read-Modify-Write *)
+empty rmw & (fre ; coe) as atomic
+
+(**********************************)
+(* Instruction execution ordering *)
+(**********************************)
+
+(* Preserved Program Order *)
+let dep = addr | data
+let rwdep = (dep | ctrl) ; [W]
+let overwrite = co | fr
+let to-w = rwdep | (overwrite & int) | (addr ; [Plain] ; wmb)
+let to-r = (addr ; [R]) | (dep ; [Marked] ; rfi)
+let ppo = to-r | to-w | (fence & int) | (po-unlock-lock-po & int)
+
+(* Propagation: Ordering from release operations and strong fences. *)
+let A-cumul(r) = (rfe ; [Marked])? ; r
+let rmw-sequence = (rf ; rmw)*
+let cumul-fence = [Marked] ; (A-cumul(strong-fence | po-rel) | wmb |
+	po-unlock-lock-po) ; [Marked] ; rmw-sequence
+let prop = [Marked] ; (overwrite & ext)? ; cumul-fence* ;
+	[Marked] ; rfe? ; [Marked]
+
+(*
+ * Happens Before: Ordering from the passage of time.
+ * No fences needed here for prop because relation confined to one process.
+ *)
+let hb = [Marked] ; (ppo | rfe | ((prop \ id) & int)) ; [Marked]
+acyclic hb as happens-before
+
+(****************************************)
+(* Write and fence propagation ordering *)
+(****************************************)
+
+(* Propagation: Each non-rf link needs a strong fence. *)
+let pb = prop ; strong-fence ; hb* ; [Marked]
+acyclic pb as propagation
+
+(*******)
+(* RCU *)
+(*******)
+
+(*
+ * Effects of read-side critical sections proceed from the rcu_read_unlock()
+ * or srcu_read_unlock() backwards on the one hand, and from the
+ * rcu_read_lock() or srcu_read_lock() forwards on the other hand.
+ *
+ * In the definition of rcu-fence below, the po term at the left-hand side
+ * of each disjunct and the po? term at the right-hand end have been factored
+ * out.  They have been moved into the definitions of rcu-link and rb.
+ * This was necessary in order to apply the "& loc" tests correctly.
+ *)
+let rcu-gp = [Sync-rcu]		(* Compare with gp *)
+let srcu-gp = [Sync-srcu]
+let rcu-rscsi = rcu-rscs^-1
+let srcu-rscsi = srcu-rscs^-1
+
+(*
+ * The synchronize_rcu() strong fence is special in that it can order not
+ * one but two non-rf relations, but only in conjunction with an RCU
+ * read-side critical section.
+ *)
+let rcu-link = po? ; hb* ; pb* ; prop ; po
+
+(*
+ * Any sequence containing at least as many grace periods as RCU read-side
+ * critical sections (joined by rcu-link) induces order like a generalized
+ * inter-CPU strong fence.
+ * Likewise for SRCU grace periods and read-side critical sections, provided
+ * the synchronize_srcu() and srcu_read_[un]lock() calls refer to the same
+ * struct srcu_struct location.
+ *)
+let rec rcu-order = rcu-gp | srcu-gp |
+	(rcu-gp ; rcu-link ; rcu-rscsi) |
+	((srcu-gp ; rcu-link ; srcu-rscsi) & loc) |
+	(rcu-rscsi ; rcu-link ; rcu-gp) |
+	((srcu-rscsi ; rcu-link ; srcu-gp) & loc) |
+	(rcu-gp ; rcu-link ; rcu-order ; rcu-link ; rcu-rscsi) |
+	((srcu-gp ; rcu-link ; rcu-order ; rcu-link ; srcu-rscsi) & loc) |
+	(rcu-rscsi ; rcu-link ; rcu-order ; rcu-link ; rcu-gp) |
+	((srcu-rscsi ; rcu-link ; rcu-order ; rcu-link ; srcu-gp) & loc) |
+	(rcu-order ; rcu-link ; rcu-order)
+let rcu-fence = po ; rcu-order ; po?
+let fence = fence | rcu-fence
+let strong-fence = strong-fence | rcu-fence
+
+(* rb orders instructions just as pb does *)
+let rb = prop ; rcu-fence ; hb* ; pb* ; [Marked]
+
+irreflexive rb as rcu
+
+(*
+ * The happens-before, propagation, and rcu constraints are all
+ * expressions of temporal ordering.  They could be replaced by
+ * a single constraint on an "executes-before" relation, xb:
+ *
+ * let xb = hb | pb | rb
+ * acyclic xb as executes-before
+ *)
+
+(*********************************)
+(* Plain accesses and data races *)
+(*********************************)
+
+(* Warn about plain writes and marked accesses in the same region *)
+let mixed-accesses = ([Plain & W] ; (po-loc \ barrier) ; [Marked]) |
+	([Marked] ; (po-loc \ barrier) ; [Plain & W])
+flag ~empty mixed-accesses as mixed-accesses
+
+(* Executes-before and visibility *)
+let xbstar = (hb | pb | rb)*
+let vis = cumul-fence* ; rfe? ; [Marked] ;
+	((strong-fence ; [Marked] ; xbstar) | (xbstar & int))
+
+(* Boundaries for lifetimes of plain accesses *)
+let w-pre-bounded = [Marked] ; (addr | fence)?
+let r-pre-bounded = [Marked] ; (addr | nonrw-fence |
+	([R4rmb] ; fencerel(Rmb) ; [~Noreturn]))?
+let w-post-bounded = fence? ; [Marked] ; rmw-sequence
+let r-post-bounded = (nonrw-fence | ([~Noreturn] ; fencerel(Rmb) ; [R4rmb]))? ;
+	[Marked]
+
+(* Visibility and executes-before for plain accesses *)
+let ww-vis = fence | (strong-fence ; xbstar ; w-pre-bounded) |
+	(w-post-bounded ; vis ; w-pre-bounded)
+let wr-vis = fence | (strong-fence ; xbstar ; r-pre-bounded) |
+	(w-post-bounded ; vis ; r-pre-bounded)
+let rw-xbstar = fence | (r-post-bounded ; xbstar ; w-pre-bounded)
+
+(* Potential races *)
+let pre-race = ext & ((Plain * M) | ((M \ IW) * Plain))
+
+(* Coherence requirements for plain accesses *)
+let wr-incoh = pre-race & rf & rw-xbstar^-1
+let rw-incoh = pre-race & fr & wr-vis^-1
+let ww-incoh = pre-race & co & ww-vis^-1
+empty (wr-incoh | rw-incoh | ww-incoh) as plain-coherence
+
+(* Actual races *)
+let ww-nonrace = ww-vis & ((Marked * W) | rw-xbstar) & ((W * Marked) | wr-vis)
+let ww-race = (pre-race & co) \ ww-nonrace
+let wr-race = (pre-race & (co? ; rf)) \ wr-vis \ rw-xbstar^-1
+let rw-race = (pre-race & fr) \ rw-xbstar
+
+flag ~empty (ww-race | wr-race | rw-race) as data-race
diff --git a/tools/memory-model/linux-kernel.cfg b/tools/memory-model/linux-kernel.cfg
new file mode 100644
index 0000000000..3c8098e99f
--- /dev/null
+++ b/tools/memory-model/linux-kernel.cfg
@@ -0,0 +1,21 @@
+macros linux-kernel.def
+bell linux-kernel.bell
+model linux-kernel.cat
+graph columns
+squished true
+showevents noregs
+movelabel true
+fontsize 8
+xscale 2.0
+yscale 1.5
+arrowsize 0.8
+showinitrf false
+showfinalrf false
+showinitwrites false
+splines spline
+pad 0.1
+edgeattr hb,color,indigo
+edgeattr co,color,blue
+edgeattr mb,color,darkgreen
+edgeattr wmb,color,darkgreen
+edgeattr rmb,color,darkgreen
diff --git a/tools/memory-model/linux-kernel.def b/tools/memory-model/linux-kernel.def
new file mode 100644
index 0000000000..88a39601f5
--- /dev/null
+++ b/tools/memory-model/linux-kernel.def
@@ -0,0 +1,119 @@
+// SPDX-License-Identifier: GPL-2.0+
+//
+// An earlier version of this file appeared in the companion webpage for
+// "Frightening small children and disconcerting grown-ups: Concurrency
+// in the Linux kernel" by Alglave, Maranget, McKenney, Parri, and Stern,
+// which appeared in ASPLOS 2018.
+
+// ONCE
+READ_ONCE(X) __load{once}(X)
+WRITE_ONCE(X,V) { __store{once}(X,V); }
+
+// Release Acquire and friends
+smp_store_release(X,V) { __store{release}(*X,V); }
+smp_load_acquire(X) __load{acquire}(*X)
+rcu_assign_pointer(X,V) { __store{release}(X,V); }
+rcu_dereference(X) __load{once}(X)
+smp_store_mb(X,V) { __store{once}(X,V); __fence{mb}; }
+
+// Fences
+smp_mb() { __fence{mb}; }
+smp_rmb() { __fence{rmb}; }
+smp_wmb() { __fence{wmb}; }
+smp_mb__before_atomic() { __fence{before-atomic}; }
+smp_mb__after_atomic() { __fence{after-atomic}; }
+smp_mb__after_spinlock() { __fence{after-spinlock}; }
+smp_mb__after_unlock_lock() { __fence{after-unlock-lock}; }
+smp_mb__after_srcu_read_unlock() { __fence{after-srcu-read-unlock}; }
+barrier() { __fence{barrier}; }
+
+// Exchange
+xchg(X,V)  __xchg{mb}(X,V)
+xchg_relaxed(X,V) __xchg{once}(X,V)
+xchg_release(X,V) __xchg{release}(X,V)
+xchg_acquire(X,V) __xchg{acquire}(X,V)
+cmpxchg(X,V,W) __cmpxchg{mb}(X,V,W)
+cmpxchg_relaxed(X,V,W) __cmpxchg{once}(X,V,W)
+cmpxchg_acquire(X,V,W) __cmpxchg{acquire}(X,V,W)
+cmpxchg_release(X,V,W) __cmpxchg{release}(X,V,W)
+
+// Spinlocks
+spin_lock(X) { __lock(X); }
+spin_unlock(X) { __unlock(X); }
+spin_trylock(X) __trylock(X)
+spin_is_locked(X) __islocked(X)
+
+// RCU
+rcu_read_lock() { __fence{rcu-lock}; }
+rcu_read_unlock() { __fence{rcu-unlock}; }
+synchronize_rcu() { __fence{sync-rcu}; }
+synchronize_rcu_expedited() { __fence{sync-rcu}; }
+
+// SRCU
+srcu_read_lock(X) __load{srcu-lock}(*X)
+srcu_read_unlock(X,Y) { __store{srcu-unlock}(*X,Y); }
+srcu_down_read(X) __load{srcu-lock}(*X)
+srcu_up_read(X,Y) { __store{srcu-unlock}(*X,Y); }
+synchronize_srcu(X)  { __srcu{sync-srcu}(X); }
+synchronize_srcu_expedited(X)  { __srcu{sync-srcu}(X); }
+
+// Atomic
+atomic_read(X) READ_ONCE(*X)
+atomic_set(X,V) { WRITE_ONCE(*X,V); }
+atomic_read_acquire(X) smp_load_acquire(X)
+atomic_set_release(X,V) { smp_store_release(X,V); }
+
+atomic_add(V,X) { __atomic_op(X,+,V); }
+atomic_sub(V,X) { __atomic_op(X,-,V); }
+atomic_inc(X)   { __atomic_op(X,+,1); }
+atomic_dec(X)   { __atomic_op(X,-,1); }
+
+atomic_add_return(V,X) __atomic_op_return{mb}(X,+,V)
+atomic_add_return_relaxed(V,X) __atomic_op_return{once}(X,+,V)
+atomic_add_return_acquire(V,X) __atomic_op_return{acquire}(X,+,V)
+atomic_add_return_release(V,X) __atomic_op_return{release}(X,+,V)
+atomic_fetch_add(V,X) __atomic_fetch_op{mb}(X,+,V)
+atomic_fetch_add_relaxed(V,X) __atomic_fetch_op{once}(X,+,V)
+atomic_fetch_add_acquire(V,X) __atomic_fetch_op{acquire}(X,+,V)
+atomic_fetch_add_release(V,X) __atomic_fetch_op{release}(X,+,V)
+
+atomic_inc_return(X) __atomic_op_return{mb}(X,+,1)
+atomic_inc_return_relaxed(X) __atomic_op_return{once}(X,+,1)
+atomic_inc_return_acquire(X) __atomic_op_return{acquire}(X,+,1)
+atomic_inc_return_release(X) __atomic_op_return{release}(X,+,1)
+atomic_fetch_inc(X) __atomic_fetch_op{mb}(X,+,1)
+atomic_fetch_inc_relaxed(X) __atomic_fetch_op{once}(X,+,1)
+atomic_fetch_inc_acquire(X) __atomic_fetch_op{acquire}(X,+,1)
+atomic_fetch_inc_release(X) __atomic_fetch_op{release}(X,+,1)
+
+atomic_sub_return(V,X) __atomic_op_return{mb}(X,-,V)
+atomic_sub_return_relaxed(V,X) __atomic_op_return{once}(X,-,V)
+atomic_sub_return_acquire(V,X) __atomic_op_return{acquire}(X,-,V)
+atomic_sub_return_release(V,X) __atomic_op_return{release}(X,-,V)
+atomic_fetch_sub(V,X) __atomic_fetch_op{mb}(X,-,V)
+atomic_fetch_sub_relaxed(V,X) __atomic_fetch_op{once}(X,-,V)
+atomic_fetch_sub_acquire(V,X) __atomic_fetch_op{acquire}(X,-,V)
+atomic_fetch_sub_release(V,X) __atomic_fetch_op{release}(X,-,V)
+
+atomic_dec_return(X) __atomic_op_return{mb}(X,-,1)
+atomic_dec_return_relaxed(X) __atomic_op_return{once}(X,-,1)
+atomic_dec_return_acquire(X) __atomic_op_return{acquire}(X,-,1)
+atomic_dec_return_release(X) __atomic_op_return{release}(X,-,1)
+atomic_fetch_dec(X) __atomic_fetch_op{mb}(X,-,1)
+atomic_fetch_dec_relaxed(X) __atomic_fetch_op{once}(X,-,1)
+atomic_fetch_dec_acquire(X) __atomic_fetch_op{acquire}(X,-,1)
+atomic_fetch_dec_release(X) __atomic_fetch_op{release}(X,-,1)
+
+atomic_xchg(X,V) __xchg{mb}(X,V)
+atomic_xchg_relaxed(X,V) __xchg{once}(X,V)
+atomic_xchg_release(X,V) __xchg{release}(X,V)
+atomic_xchg_acquire(X,V) __xchg{acquire}(X,V)
+atomic_cmpxchg(X,V,W) __cmpxchg{mb}(X,V,W)
+atomic_cmpxchg_relaxed(X,V,W) __cmpxchg{once}(X,V,W)
+atomic_cmpxchg_acquire(X,V,W) __cmpxchg{acquire}(X,V,W)
+atomic_cmpxchg_release(X,V,W) __cmpxchg{release}(X,V,W)
+
+atomic_sub_and_test(V,X) __atomic_op_return{mb}(X,-,V) == 0
+atomic_dec_and_test(X)  __atomic_op_return{mb}(X,-,1) == 0
+atomic_inc_and_test(X)  __atomic_op_return{mb}(X,+,1) == 0
+atomic_add_negative(V,X) __atomic_op_return{mb}(X,+,V) < 0
diff --git a/tools/memory-model/litmus-tests/.gitignore b/tools/memory-model/litmus-tests/.gitignore
new file mode 100644
index 0000000000..19c379cf06
--- /dev/null
+++ b/tools/memory-model/litmus-tests/.gitignore
@@ -0,0 +1,2 @@
+# SPDX-License-Identifier: GPL-2.0-only
+*.litmus.*
diff --git a/tools/memory-model/litmus-tests/CoRR+poonceonce+Once.litmus b/tools/memory-model/litmus-tests/CoRR+poonceonce+Once.litmus
new file mode 100644
index 0000000000..967f9f2a62
--- /dev/null
+++ b/tools/memory-model/litmus-tests/CoRR+poonceonce+Once.litmus
@@ -0,0 +1,26 @@
+C CoRR+poonceonce+Once
+
+(*
+ * Result: Never
+ *
+ * Test of read-read coherence, that is, whether or not two successive
+ * reads from the same variable are ordered.
+ *)
+
+{}
+
+P0(int *x)
+{
+	WRITE_ONCE(*x, 1);
+}
+
+P1(int *x)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*x);
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 1:r1=0)
diff --git a/tools/memory-model/litmus-tests/CoRW+poonceonce+Once.litmus b/tools/memory-model/litmus-tests/CoRW+poonceonce+Once.litmus
new file mode 100644
index 0000000000..4635739f39
--- /dev/null
+++ b/tools/memory-model/litmus-tests/CoRW+poonceonce+Once.litmus
@@ -0,0 +1,25 @@
+C CoRW+poonceonce+Once
+
+(*
+ * Result: Never
+ *
+ * Test of read-write coherence, that is, whether or not a read from
+ * a given variable and a later write to that same variable are ordered.
+ *)
+
+{}
+
+P0(int *x)
+{
+	int r0;
+
+	r0 = READ_ONCE(*x);
+	WRITE_ONCE(*x, 1);
+}
+
+P1(int *x)
+{
+	WRITE_ONCE(*x, 2);
+}
+
+exists (x=2 /\ 0:r0=2)
diff --git a/tools/memory-model/litmus-tests/CoWR+poonceonce+Once.litmus b/tools/memory-model/litmus-tests/CoWR+poonceonce+Once.litmus
new file mode 100644
index 0000000000..bb068c92d8
--- /dev/null
+++ b/tools/memory-model/litmus-tests/CoWR+poonceonce+Once.litmus
@@ -0,0 +1,25 @@
+C CoWR+poonceonce+Once
+
+(*
+ * Result: Never
+ *
+ * Test of write-read coherence, that is, whether or not a write to a
+ * given variable and a later read from that same variable are ordered.
+ *)
+
+{}
+
+P0(int *x)
+{
+	int r0;
+
+	WRITE_ONCE(*x, 1);
+	r0 = READ_ONCE(*x);
+}
+
+P1(int *x)
+{
+	WRITE_ONCE(*x, 2);
+}
+
+exists (x=1 /\ 0:r0=2)
diff --git a/tools/memory-model/litmus-tests/CoWW+poonceonce.litmus b/tools/memory-model/litmus-tests/CoWW+poonceonce.litmus
new file mode 100644
index 0000000000..0d9f0a9587
--- /dev/null
+++ b/tools/memory-model/litmus-tests/CoWW+poonceonce.litmus
@@ -0,0 +1,18 @@
+C CoWW+poonceonce
+
+(*
+ * Result: Never
+ *
+ * Test of write-write coherence, that is, whether or not two successive
+ * writes to the same variable are ordered.
+ *)
+
+{}
+
+P0(int *x)
+{
+	WRITE_ONCE(*x, 1);
+	WRITE_ONCE(*x, 2);
+}
+
+exists (x=1)
diff --git a/tools/memory-model/litmus-tests/IRIW+fencembonceonces+OnceOnce.litmus b/tools/memory-model/litmus-tests/IRIW+fencembonceonces+OnceOnce.litmus
new file mode 100644
index 0000000000..e729d2776e
--- /dev/null
+++ b/tools/memory-model/litmus-tests/IRIW+fencembonceonces+OnceOnce.litmus
@@ -0,0 +1,45 @@
+C IRIW+fencembonceonces+OnceOnce
+
+(*
+ * Result: Never
+ *
+ * Test of independent reads from independent writes with smp_mb()
+ * between each pairs of reads.  In other words, is smp_mb() sufficient to
+ * cause two different reading processes to agree on the order of a pair
+ * of writes, where each write is to a different variable by a different
+ * process?  This litmus test exercises LKMM's "propagation" rule.
+ *)
+
+{}
+
+P0(int *x)
+{
+	WRITE_ONCE(*x, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*x);
+	smp_mb();
+	r1 = READ_ONCE(*y);
+}
+
+P2(int *y)
+{
+	WRITE_ONCE(*y, 1);
+}
+
+P3(int *x, int *y)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*y);
+	smp_mb();
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 1:r1=0 /\ 3:r0=1 /\ 3:r1=0)
diff --git a/tools/memory-model/litmus-tests/IRIW+poonceonces+OnceOnce.litmus b/tools/memory-model/litmus-tests/IRIW+poonceonces+OnceOnce.litmus
new file mode 100644
index 0000000000..4b54dd6a6c
--- /dev/null
+++ b/tools/memory-model/litmus-tests/IRIW+poonceonces+OnceOnce.litmus
@@ -0,0 +1,43 @@
+C IRIW+poonceonces+OnceOnce
+
+(*
+ * Result: Sometimes
+ *
+ * Test of independent reads from independent writes with nothing
+ * between each pairs of reads.  In other words, is anything at all
+ * needed to cause two different reading processes to agree on the order
+ * of a pair of writes, where each write is to a different variable by a
+ * different process?
+ *)
+
+{}
+
+P0(int *x)
+{
+	WRITE_ONCE(*x, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*x);
+	r1 = READ_ONCE(*y);
+}
+
+P2(int *y)
+{
+	WRITE_ONCE(*y, 1);
+}
+
+P3(int *x, int *y)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*y);
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 1:r1=0 /\ 3:r0=1 /\ 3:r1=0)
diff --git a/tools/memory-model/litmus-tests/ISA2+pooncelock+pooncelock+pombonce.litmus b/tools/memory-model/litmus-tests/ISA2+pooncelock+pooncelock+pombonce.litmus
new file mode 100644
index 0000000000..094d58df77
--- /dev/null
+++ b/tools/memory-model/litmus-tests/ISA2+pooncelock+pooncelock+pombonce.litmus
@@ -0,0 +1,40 @@
+C ISA2+pooncelock+pooncelock+pombonce
+
+(*
+ * Result: Never
+ *
+ * This test shows that write-write ordering provided by locks
+ * (in P0() and P1()) is visible to external process P2().
+ *)
+
+{}
+
+P0(int *x, int *y, spinlock_t *mylock)
+{
+	spin_lock(mylock);
+	WRITE_ONCE(*x, 1);
+	WRITE_ONCE(*y, 1);
+	spin_unlock(mylock);
+}
+
+P1(int *y, int *z, spinlock_t *mylock)
+{
+	int r0;
+
+	spin_lock(mylock);
+	r0 = READ_ONCE(*y);
+	WRITE_ONCE(*z, 1);
+	spin_unlock(mylock);
+}
+
+P2(int *x, int *z)
+{
+	int r1;
+	int r2;
+
+	r2 = READ_ONCE(*z);
+	smp_mb();
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 2:r2=1 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/ISA2+poonceonces.litmus b/tools/memory-model/litmus-tests/ISA2+poonceonces.litmus
new file mode 100644
index 0000000000..b321aa6f4e
--- /dev/null
+++ b/tools/memory-model/litmus-tests/ISA2+poonceonces.litmus
@@ -0,0 +1,37 @@
+C ISA2+poonceonces
+
+(*
+ * Result: Sometimes
+ *
+ * Given a release-acquire chain ordering the first process's store
+ * against the last process's load, is ordering preserved if all of the
+ * smp_store_release() invocations are replaced by WRITE_ONCE() and all
+ * of the smp_load_acquire() invocations are replaced by READ_ONCE()?
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	WRITE_ONCE(*x, 1);
+	WRITE_ONCE(*y, 1);
+}
+
+P1(int *y, int *z)
+{
+	int r0;
+
+	r0 = READ_ONCE(*y);
+	WRITE_ONCE(*z, 1);
+}
+
+P2(int *x, int *z)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*z);
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 2:r0=1 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/ISA2+pooncerelease+poacquirerelease+poacquireonce.litmus b/tools/memory-model/litmus-tests/ISA2+pooncerelease+poacquirerelease+poacquireonce.litmus
new file mode 100644
index 0000000000..025b0462ec
--- /dev/null
+++ b/tools/memory-model/litmus-tests/ISA2+pooncerelease+poacquirerelease+poacquireonce.litmus
@@ -0,0 +1,39 @@
+C ISA2+pooncerelease+poacquirerelease+poacquireonce
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates that a release-acquire chain suffices
+ * to order P0()'s initial write against P2()'s final read.  The reason
+ * that the release-acquire chain suffices is because in all but one
+ * case (P2() to P0()), each process reads from the preceding process's
+ * write.  In memory-model-speak, there is only one non-reads-from
+ * (AKA non-rf) link, so release-acquire is all that is needed.
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	WRITE_ONCE(*x, 1);
+	smp_store_release(y, 1);
+}
+
+P1(int *y, int *z)
+{
+	int r0;
+
+	r0 = smp_load_acquire(y);
+	smp_store_release(z, 1);
+}
+
+P2(int *x, int *z)
+{
+	int r0;
+	int r1;
+
+	r0 = smp_load_acquire(z);
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 2:r0=1 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/LB+fencembonceonce+ctrlonceonce.litmus b/tools/memory-model/litmus-tests/LB+fencembonceonce+ctrlonceonce.litmus
new file mode 100644
index 0000000000..4727f5aaf0
--- /dev/null
+++ b/tools/memory-model/litmus-tests/LB+fencembonceonce+ctrlonceonce.litmus
@@ -0,0 +1,34 @@
+C LB+fencembonceonce+ctrlonceonce
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates that lightweight ordering suffices for
+ * the load-buffering pattern, in other words, preventing all processes
+ * reading from the preceding process's write.  In this example, the
+ * combination of a control dependency and a full memory barrier are enough
+ * to do the trick.  (But the full memory barrier could be replaced with
+ * another control dependency and order would still be maintained.)
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*x);
+	if (r0)
+		WRITE_ONCE(*y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*y);
+	smp_mb();
+	WRITE_ONCE(*x, 1);
+}
+
+exists (0:r0=1 /\ 1:r0=1)
diff --git a/tools/memory-model/litmus-tests/LB+poacquireonce+pooncerelease.litmus b/tools/memory-model/litmus-tests/LB+poacquireonce+pooncerelease.litmus
new file mode 100644
index 0000000000..07b9904b0e
--- /dev/null
+++ b/tools/memory-model/litmus-tests/LB+poacquireonce+pooncerelease.litmus
@@ -0,0 +1,29 @@
+C LB+poacquireonce+pooncerelease
+
+(*
+ * Result: Never
+ *
+ * Does a release-acquire pair suffice for the load-buffering litmus
+ * test, where each process reads from one of two variables then writes
+ * to the other?
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*x);
+	smp_store_release(y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	r0 = smp_load_acquire(y);
+	WRITE_ONCE(*x, 1);
+}
+
+exists (0:r0=1 /\ 1:r0=1)
diff --git a/tools/memory-model/litmus-tests/LB+poonceonces.litmus b/tools/memory-model/litmus-tests/LB+poonceonces.litmus
new file mode 100644
index 0000000000..74c49cb3c3
--- /dev/null
+++ b/tools/memory-model/litmus-tests/LB+poonceonces.litmus
@@ -0,0 +1,28 @@
+C LB+poonceonces
+
+(*
+ * Result: Sometimes
+ *
+ * Can the counter-intuitive outcome for the load-buffering pattern
+ * be prevented even with no explicit ordering?
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*x);
+	WRITE_ONCE(*y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*y);
+	WRITE_ONCE(*x, 1);
+}
+
+exists (0:r0=1 /\ 1:r0=1)
diff --git a/tools/memory-model/litmus-tests/LB+unlocklockonceonce+poacquireonce.litmus b/tools/memory-model/litmus-tests/LB+unlocklockonceonce+poacquireonce.litmus
new file mode 100644
index 0000000000..eb34123a6f
--- /dev/null
+++ b/tools/memory-model/litmus-tests/LB+unlocklockonceonce+poacquireonce.litmus
@@ -0,0 +1,35 @@
+C LB+unlocklockonceonce+poacquireonce
+
+(*
+ * Result: Never
+ *
+ * If two locked critical sections execute on the same CPU, all accesses
+ * in the first must execute before any accesses in the second, even if the
+ * critical sections are protected by different locks.  Note: Even when a
+ * write executes before a read, their memory effects can be reordered from
+ * the viewpoint of another CPU (the kind of reordering allowed by TSO).
+ *)
+
+{}
+
+P0(spinlock_t *s, spinlock_t *t, int *x, int *y)
+{
+	int r1;
+
+	spin_lock(s);
+	r1 = READ_ONCE(*x);
+	spin_unlock(s);
+	spin_lock(t);
+	WRITE_ONCE(*y, 1);
+	spin_unlock(t);
+}
+
+P1(int *x, int *y)
+{
+	int r2;
+
+	r2 = smp_load_acquire(y);
+	WRITE_ONCE(*x, 1);
+}
+
+exists (0:r1=1 /\ 1:r2=1)
diff --git a/tools/memory-model/litmus-tests/MP+fencewmbonceonce+fencermbonceonce.litmus b/tools/memory-model/litmus-tests/MP+fencewmbonceonce+fencermbonceonce.litmus
new file mode 100644
index 0000000000..f8ca122985
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+fencewmbonceonce+fencermbonceonce.litmus
@@ -0,0 +1,30 @@
+C MP+fencewmbonceonce+fencermbonceonce
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates that smp_wmb() and smp_rmb() provide
+ * sufficient ordering for the message-passing pattern.  However, it
+ * is usually better to use smp_store_release() and smp_load_acquire().
+ *)
+
+{}
+
+P0(int *buf, int *flag) // Producer
+{
+	WRITE_ONCE(*buf, 1);
+	smp_wmb();
+	WRITE_ONCE(*flag, 1);
+}
+
+P1(int *buf, int *flag) // Consumer
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*flag);
+	smp_rmb();
+	r1 = READ_ONCE(*buf);
+}
+
+exists (1:r0=1 /\ 1:r1=0) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+onceassign+derefonce.litmus b/tools/memory-model/litmus-tests/MP+onceassign+derefonce.litmus
new file mode 100644
index 0000000000..d84160b9c1
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+onceassign+derefonce.litmus
@@ -0,0 +1,33 @@
+C MP+onceassign+derefonce
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates that rcu_assign_pointer() and
+ * rcu_dereference() suffice to ensure that an RCU reader will not see
+ * pre-initialization garbage when it traverses an RCU-protected data
+ * structure containing a newly inserted element.
+ *)
+
+{
+p=y;
+}
+
+P0(int *x, int **p) // Producer
+{
+	WRITE_ONCE(*x, 1);
+	rcu_assign_pointer(*p, x);
+}
+
+P1(int *x, int **p) // Consumer
+{
+	int *r0;
+	int r1;
+
+	rcu_read_lock();
+	r0 = rcu_dereference(*p);
+	r1 = READ_ONCE(*r0);
+	rcu_read_unlock();
+}
+
+exists (1:r0=x /\ 1:r1=0) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+polockmbonce+poacquiresilsil.litmus b/tools/memory-model/litmus-tests/MP+polockmbonce+poacquiresilsil.litmus
new file mode 100644
index 0000000000..ba91cc63e1
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+polockmbonce+poacquiresilsil.litmus
@@ -0,0 +1,34 @@
+C MP+polockmbonce+poacquiresilsil
+
+(*
+ * Result: Never
+ *
+ * Do spinlocks combined with smp_mb__after_spinlock() provide order
+ * to outside observers using spin_is_locked() to sense the lock-held
+ * state, ordered by acquire?  Note that when the first spin_is_locked()
+ * returns false and the second true, we know that the smp_load_acquire()
+ * executed before the lock was acquired (loosely speaking).
+ *)
+
+{}
+
+P0(spinlock_t *lo, int *x) // Producer
+{
+	spin_lock(lo);
+	smp_mb__after_spinlock();
+	WRITE_ONCE(*x, 1);
+	spin_unlock(lo);
+}
+
+P1(spinlock_t *lo, int *x) // Consumer
+{
+	int r1;
+	int r2;
+	int r3;
+
+	r1 = smp_load_acquire(x);
+	r2 = spin_is_locked(lo);
+	r3 = spin_is_locked(lo);
+}
+
+exists (1:r1=1 /\ 1:r2=0 /\ 1:r3=1) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+polockonce+poacquiresilsil.litmus b/tools/memory-model/litmus-tests/MP+polockonce+poacquiresilsil.litmus
new file mode 100644
index 0000000000..a5ea3ed8f5
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+polockonce+poacquiresilsil.litmus
@@ -0,0 +1,33 @@
+C MP+polockonce+poacquiresilsil
+
+(*
+ * Result: Sometimes
+ *
+ * Do spinlocks provide order to outside observers using spin_is_locked()
+ * to sense the lock-held state, ordered by acquire?  Note that when the
+ * first spin_is_locked() returns false and the second true, we know that
+ * the smp_load_acquire() executed before the lock was acquired (loosely
+ * speaking).
+ *)
+
+{}
+
+P0(spinlock_t *lo, int *x) // Producer
+{
+	spin_lock(lo);
+	WRITE_ONCE(*x, 1);
+	spin_unlock(lo);
+}
+
+P1(spinlock_t *lo, int *x) // Consumer
+{
+	int r1;
+	int r2;
+	int r3;
+
+	r1 = smp_load_acquire(x);
+	r2 = spin_is_locked(lo);
+	r3 = spin_is_locked(lo);
+}
+
+exists (1:r1=1 /\ 1:r2=0 /\ 1:r3=1) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+polocks.litmus b/tools/memory-model/litmus-tests/MP+polocks.litmus
new file mode 100644
index 0000000000..e6af05f700
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+polocks.litmus
@@ -0,0 +1,35 @@
+C MP+polocks
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates how lock acquisitions and releases can
+ * stand in for smp_load_acquire() and smp_store_release(), respectively.
+ * In other words, when holding a given lock (or indeed after releasing a
+ * given lock), a CPU is not only guaranteed to see the accesses that other
+ * CPUs made while previously holding that lock, it is also guaranteed
+ * to see all prior accesses by those other CPUs.
+ *)
+
+{}
+
+P0(int *buf, int *flag, spinlock_t *mylock) // Producer
+{
+	WRITE_ONCE(*buf, 1);
+	spin_lock(mylock);
+	WRITE_ONCE(*flag, 1);
+	spin_unlock(mylock);
+}
+
+P1(int *buf, int *flag, spinlock_t *mylock) // Consumer
+{
+	int r0;
+	int r1;
+
+	spin_lock(mylock);
+	r0 = READ_ONCE(*flag);
+	spin_unlock(mylock);
+	r1 = READ_ONCE(*buf);
+}
+
+exists (1:r0=1 /\ 1:r1=0) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+poonceonces.litmus b/tools/memory-model/litmus-tests/MP+poonceonces.litmus
new file mode 100644
index 0000000000..ba9c99c6cf
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+poonceonces.litmus
@@ -0,0 +1,27 @@
+C MP+poonceonces
+
+(*
+ * Result: Sometimes
+ *
+ * Can the counter-intuitive message-passing outcome be prevented with
+ * no ordering at all?
+ *)
+
+{}
+
+P0(int *buf, int *flag) // Producer
+{
+	WRITE_ONCE(*buf, 1);
+	WRITE_ONCE(*flag, 1);
+}
+
+P1(int *buf, int *flag) // Consumer
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*flag);
+	r1 = READ_ONCE(*buf);
+}
+
+exists (1:r0=1 /\ 1:r1=0) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+pooncerelease+poacquireonce.litmus b/tools/memory-model/litmus-tests/MP+pooncerelease+poacquireonce.litmus
new file mode 100644
index 0000000000..f174bfe617
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+pooncerelease+poacquireonce.litmus
@@ -0,0 +1,28 @@
+C MP+pooncerelease+poacquireonce
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates that smp_store_release() and
+ * smp_load_acquire() provide sufficient ordering for the message-passing
+ * pattern.
+ *)
+
+{}
+
+P0(int *buf, int *flag) // Producer
+{
+	WRITE_ONCE(*buf, 1);
+	smp_store_release(flag, 1);
+}
+
+P1(int *buf, int *flag) // Consumer
+{
+	int r0;
+	int r1;
+
+	r0 = smp_load_acquire(flag);
+	r1 = READ_ONCE(*buf);
+}
+
+exists (1:r0=1 /\ 1:r1=0) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+porevlocks.litmus b/tools/memory-model/litmus-tests/MP+porevlocks.litmus
new file mode 100644
index 0000000000..b959914116
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+porevlocks.litmus
@@ -0,0 +1,35 @@
+C MP+porevlocks
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates how lock acquisitions and releases can
+ * stand in for smp_load_acquire() and smp_store_release(), respectively.
+ * In other words, when holding a given lock (or indeed after releasing a
+ * given lock), a CPU is not only guaranteed to see the accesses that other
+ * CPUs made while previously holding that lock, it is also guaranteed to
+ * see all prior accesses by those other CPUs.
+ *)
+
+{}
+
+P0(int *buf, int *flag, spinlock_t *mylock) // Consumer
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*flag);
+	spin_lock(mylock);
+	r1 = READ_ONCE(*buf);
+	spin_unlock(mylock);
+}
+
+P1(int *buf, int *flag, spinlock_t *mylock) // Producer
+{
+	spin_lock(mylock);
+	WRITE_ONCE(*buf, 1);
+	spin_unlock(mylock);
+	WRITE_ONCE(*flag, 1);
+}
+
+exists (0:r0=1 /\ 0:r1=0) (* Bad outcome. *)
diff --git a/tools/memory-model/litmus-tests/MP+unlocklockonceonce+fencermbonceonce.litmus b/tools/memory-model/litmus-tests/MP+unlocklockonceonce+fencermbonceonce.litmus
new file mode 100644
index 0000000000..2feb1398be
--- /dev/null
+++ b/tools/memory-model/litmus-tests/MP+unlocklockonceonce+fencermbonceonce.litmus
@@ -0,0 +1,33 @@
+C MP+unlocklockonceonce+fencermbonceonce
+
+(*
+ * Result: Never
+ *
+ * If two locked critical sections execute on the same CPU, stores in the
+ * first must propagate to each CPU before stores in the second do, even if
+ * the critical sections are protected by different locks.
+ *)
+
+{}
+
+P0(spinlock_t *s, spinlock_t *t, int *x, int *y)
+{
+	spin_lock(s);
+	WRITE_ONCE(*x, 1);
+	spin_unlock(s);
+	spin_lock(t);
+	WRITE_ONCE(*y, 1);
+	spin_unlock(t);
+}
+
+P1(int *x, int *y)
+{
+	int r1;
+	int r2;
+
+	r1 = READ_ONCE(*y);
+	smp_rmb();
+	r2 = READ_ONCE(*x);
+}
+
+exists (1:r1=1 /\ 1:r2=0)
diff --git a/tools/memory-model/litmus-tests/R+fencembonceonces.litmus b/tools/memory-model/litmus-tests/R+fencembonceonces.litmus
new file mode 100644
index 0000000000..222a0b850b
--- /dev/null
+++ b/tools/memory-model/litmus-tests/R+fencembonceonces.litmus
@@ -0,0 +1,30 @@
+C R+fencembonceonces
+
+(*
+ * Result: Never
+ *
+ * This is the fully ordered (via smp_mb()) version of one of the classic
+ * counterintuitive litmus tests that illustrates the effects of store
+ * propagation delays.  Note that weakening either of the barriers would
+ * cause the resulting test to be allowed.
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	WRITE_ONCE(*x, 1);
+	smp_mb();
+	WRITE_ONCE(*y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	WRITE_ONCE(*y, 2);
+	smp_mb();
+	r0 = READ_ONCE(*x);
+}
+
+exists (y=2 /\ 1:r0=0)
diff --git a/tools/memory-model/litmus-tests/R+poonceonces.litmus b/tools/memory-model/litmus-tests/R+poonceonces.litmus
new file mode 100644
index 0000000000..5386f128a1
--- /dev/null
+++ b/tools/memory-model/litmus-tests/R+poonceonces.litmus
@@ -0,0 +1,27 @@
+C R+poonceonces
+
+(*
+ * Result: Sometimes
+ *
+ * This is the unordered (thus lacking smp_mb()) version of one of the
+ * classic counterintuitive litmus tests that illustrates the effects of
+ * store propagation delays.
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	WRITE_ONCE(*x, 1);
+	WRITE_ONCE(*y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	WRITE_ONCE(*y, 2);
+	r0 = READ_ONCE(*x);
+}
+
+exists (y=2 /\ 1:r0=0)
diff --git a/tools/memory-model/litmus-tests/README b/tools/memory-model/litmus-tests/README
new file mode 100644
index 0000000000..d311a0ff1a
--- /dev/null
+++ b/tools/memory-model/litmus-tests/README
@@ -0,0 +1,261 @@
+============
+LITMUS TESTS
+============
+
+CoRR+poonceonce+Once.litmus
+	Test of read-read coherence, that is, whether or not two
+	successive reads from the same variable are ordered.
+
+CoRW+poonceonce+Once.litmus
+	Test of read-write coherence, that is, whether or not a read
+	from a given variable followed by a write to that same variable
+	are ordered.
+
+CoWR+poonceonce+Once.litmus
+	Test of write-read coherence, that is, whether or not a write
+	to a given variable followed by a read from that same variable
+	are ordered.
+
+CoWW+poonceonce.litmus
+	Test of write-write coherence, that is, whether or not two
+	successive writes to the same variable are ordered.
+
+IRIW+fencembonceonces+OnceOnce.litmus
+	Test of independent reads from independent writes with smp_mb()
+	between each pairs of reads.  In other words, is smp_mb()
+	sufficient to cause two different reading processes to agree on
+	the order of a pair of writes, where each write is to a different
+	variable by a different process?  This litmus test is forbidden
+	by LKMM's propagation rule.
+
+IRIW+poonceonces+OnceOnce.litmus
+	Test of independent reads from independent writes with nothing
+	between each pairs of reads.  In other words, is anything at all
+	needed to cause two different reading processes to agree on the
+	order of a pair of writes, where each write is to a different
+	variable by a different process?
+
+ISA2+pooncelock+pooncelock+pombonce.litmus
+	Tests whether the ordering provided by a lock-protected S
+	litmus test is visible to an external process whose accesses are
+	separated by smp_mb().  This addition of an external process to
+	S is otherwise known as ISA2.
+
+ISA2+poonceonces.litmus
+	As below, but with store-release replaced with WRITE_ONCE()
+	and load-acquire replaced with READ_ONCE().
+
+ISA2+pooncerelease+poacquirerelease+poacquireonce.litmus
+	Can a release-acquire chain order a prior store against
+	a later load?
+
+LB+fencembonceonce+ctrlonceonce.litmus
+	Does a control dependency and an smp_mb() suffice for the
+	load-buffering litmus test, where each process reads from one
+	of two variables then writes to the other?
+
+LB+poacquireonce+pooncerelease.litmus
+	Does a release-acquire pair suffice for the load-buffering
+	litmus test, where each process reads from one of two variables then
+	writes to the other?
+
+LB+poonceonces.litmus
+	As above, but with store-release replaced with WRITE_ONCE()
+	and load-acquire replaced with READ_ONCE().
+
+LB+unlocklockonceonce+poacquireonce.litmus
+	Does a unlock+lock pair provides ordering guarantee between a
+	load and a store?
+
+MP+onceassign+derefonce.litmus
+	As below, but with rcu_assign_pointer() and an rcu_dereference().
+
+MP+polockmbonce+poacquiresilsil.litmus
+	Protect the access with a lock and an smp_mb__after_spinlock()
+	in one process, and use an acquire load followed by a pair of
+	spin_is_locked() calls in the other process.
+
+MP+polockonce+poacquiresilsil.litmus
+	Protect the access with a lock in one process, and use an
+	acquire load followed by a pair of spin_is_locked() calls
+	in the other process.
+
+MP+polocks.litmus
+	As below, but with the second access of the writer process
+	and the first access of reader process protected by a lock.
+
+MP+poonceonces.litmus
+	As below, but without the smp_rmb() and smp_wmb().
+
+MP+pooncerelease+poacquireonce.litmus
+	As below, but with a release-acquire chain.
+
+MP+porevlocks.litmus
+	As below, but with the first access of the writer process
+	and the second access of reader process protected by a lock.
+
+MP+unlocklockonceonce+fencermbonceonce.litmus
+	Does a unlock+lock pair provides ordering guarantee between a
+	store and another store?
+
+MP+fencewmbonceonce+fencermbonceonce.litmus
+	Does a smp_wmb() (between the stores) and an smp_rmb() (between
+	the loads) suffice for the message-passing litmus test, where one
+	process writes data and then a flag, and the other process reads
+	the flag and then the data.  (This is similar to the ISA2 tests,
+	but with two processes instead of three.)
+
+R+fencembonceonces.litmus
+	This is the fully ordered (via smp_mb()) version of one of
+	the classic counterintuitive litmus tests that illustrates the
+	effects of store propagation delays.
+
+R+poonceonces.litmus
+	As above, but without the smp_mb() invocations.
+
+SB+fencembonceonces.litmus
+	This is the fully ordered (again, via smp_mb() version of store
+	buffering, which forms the core of Dekker's mutual-exclusion
+	algorithm.
+
+SB+poonceonces.litmus
+	As above, but without the smp_mb() invocations.
+
+SB+rfionceonce-poonceonces.litmus
+	This litmus test demonstrates that LKMM is not fully multicopy
+	atomic.  (Neither is it other multicopy atomic.)  This litmus test
+	also demonstrates the "locations" debugging aid, which designates
+	additional registers and locations to be printed out in the dump
+	of final states in the herd7 output.  Without the "locations"
+	statement, only those registers and locations mentioned in the
+	"exists" clause will be printed.
+
+S+poonceonces.litmus
+	As below, but without the smp_wmb() and acquire load.
+
+S+fencewmbonceonce+poacquireonce.litmus
+	Can a smp_wmb(), instead of a release, and an acquire order
+	a prior store against a subsequent store?
+
+WRC+poonceonces+Once.litmus
+WRC+pooncerelease+fencermbonceonce+Once.litmus
+	These two are members of an extension of the MP litmus-test
+	class in which the first write is moved to a separate process.
+	The second is forbidden because smp_store_release() is
+	A-cumulative in LKMM.
+
+Z6.0+pooncelock+pooncelock+pombonce.litmus
+	Is the ordering provided by a spin_unlock() and a subsequent
+	spin_lock() sufficient to make ordering apparent to accesses
+	by a process not holding the lock?
+
+Z6.0+pooncelock+poonceLock+pombonce.litmus
+	As above, but with smp_mb__after_spinlock() immediately
+	following the spin_lock().
+
+Z6.0+pooncerelease+poacquirerelease+fencembonceonce.litmus
+	Is the ordering provided by a release-acquire chain sufficient
+	to make ordering apparent to accesses by a process that does
+	not participate in that release-acquire chain?
+
+A great many more litmus tests are available here:
+
+	https://github.com/paulmckrcu/litmus
+
+==================
+LITMUS TEST NAMING
+==================
+
+Litmus tests are usually named based on their contents, which means that
+looking at the name tells you what the litmus test does.  The naming
+scheme covers litmus tests having a single cycle that passes through
+each process exactly once, so litmus tests not fitting this description
+are named on an ad-hoc basis.
+
+The structure of a litmus-test name is the litmus-test class, a plus
+sign ("+"), and one string for each process, separated by plus signs.
+The end of the name is ".litmus".
+
+The litmus-test classes may be found in the infamous test6.pdf:
+https://www.cl.cam.ac.uk/~pes20/ppc-supplemental/test6.pdf
+Each class defines the pattern of accesses and of the variables accessed.
+For example, if the one process writes to a pair of variables, and
+the other process reads from these same variables, the corresponding
+litmus-test class is "MP" (message passing), which may be found on the
+left-hand end of the second row of tests on page one of test6.pdf.
+
+The strings used to identify the actions carried out by each process are
+complex due to a desire to have short(er) names.  Thus, there is a tool to
+generate these strings from a given litmus test's actions.  For example,
+consider the processes from SB+rfionceonce-poonceonces.litmus:
+
+	P0(int *x, int *y)
+	{
+		int r1;
+		int r2;
+
+		WRITE_ONCE(*x, 1);
+		r1 = READ_ONCE(*x);
+		r2 = READ_ONCE(*y);
+	}
+
+	P1(int *x, int *y)
+	{
+		int r3;
+		int r4;
+
+		WRITE_ONCE(*y, 1);
+		r3 = READ_ONCE(*y);
+		r4 = READ_ONCE(*x);
+	}
+
+The next step is to construct a space-separated list of descriptors,
+interleaving descriptions of the relation between a pair of consecutive
+accesses with descriptions of the second access in the pair.
+
+P0()'s WRITE_ONCE() is read by its first READ_ONCE(), which is a
+reads-from link (rf) and internal to the P0() process.  This is
+"rfi", which is an abbreviation for "reads-from internal".  Because
+some of the tools string these abbreviations together with space
+characters separating processes, the first character is capitalized,
+resulting in "Rfi".
+
+P0()'s second access is a READ_ONCE(), as opposed to (for example)
+smp_load_acquire(), so next is "Once".  Thus far, we have "Rfi Once".
+
+P0()'s third access is also a READ_ONCE(), but to y rather than x.
+This is related to P0()'s second access by program order ("po"),
+to a different variable ("d"), and both accesses are reads ("RR").
+The resulting descriptor is "PodRR".  Because P0()'s third access is
+READ_ONCE(), we add another "Once" descriptor.
+
+A from-read ("fre") relation links P0()'s third to P1()'s first
+access, and the resulting descriptor is "Fre".  P1()'s first access is
+WRITE_ONCE(), which as before gives the descriptor "Once".  The string
+thus far is thus "Rfi Once PodRR Once Fre Once".
+
+The remainder of P1() is similar to P0(), which means we add
+"Rfi Once PodRR Once".  Another fre links P1()'s last access to
+P0()'s first access, which is WRITE_ONCE(), so we add "Fre Once".
+The full string is thus:
+
+	Rfi Once PodRR Once Fre Once Rfi Once PodRR Once Fre Once
+
+This string can be given to the "norm7" and "classify7" tools to
+produce the name:
+
+	$ norm7 -bell linux-kernel.bell \
+		Rfi Once PodRR Once Fre Once Rfi Once PodRR Once Fre Once | \
+	  sed -e 's/:.*//g'
+	SB+rfionceonce-poonceonces
+
+Adding the ".litmus" suffix: SB+rfionceonce-poonceonces.litmus
+
+The descriptors that describe connections between consecutive accesses
+within the cycle through a given litmus test can be provided by the herd7
+tool (Rfi, Po, Fre, and so on) or by the linux-kernel.bell file (Once,
+Release, Acquire, and so on).
+
+To see the full list of descriptors, execute the following command:
+
+	$ diyone7 -bell linux-kernel.bell -show edges
diff --git a/tools/memory-model/litmus-tests/S+fencewmbonceonce+poacquireonce.litmus b/tools/memory-model/litmus-tests/S+fencewmbonceonce+poacquireonce.litmus
new file mode 100644
index 0000000000..18479823cd
--- /dev/null
+++ b/tools/memory-model/litmus-tests/S+fencewmbonceonce+poacquireonce.litmus
@@ -0,0 +1,27 @@
+C S+fencewmbonceonce+poacquireonce
+
+(*
+ * Result: Never
+ *
+ * Can a smp_wmb(), instead of a release, and an acquire order a prior
+ * store against a subsequent store?
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	WRITE_ONCE(*x, 2);
+	smp_wmb();
+	WRITE_ONCE(*y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	r0 = smp_load_acquire(y);
+	WRITE_ONCE(*x, 1);
+}
+
+exists (x=2 /\ 1:r0=1)
diff --git a/tools/memory-model/litmus-tests/S+poonceonces.litmus b/tools/memory-model/litmus-tests/S+poonceonces.litmus
new file mode 100644
index 0000000000..8c9c2f81a5
--- /dev/null
+++ b/tools/memory-model/litmus-tests/S+poonceonces.litmus
@@ -0,0 +1,28 @@
+C S+poonceonces
+
+(*
+ * Result: Sometimes
+ *
+ * Starting with a two-process release-acquire chain ordering P0()'s
+ * first store against P1()'s final load, if the smp_store_release()
+ * is replaced by WRITE_ONCE() and the smp_load_acquire() replaced by
+ * READ_ONCE(), is ordering preserved?
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	WRITE_ONCE(*x, 2);
+	WRITE_ONCE(*y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*y);
+	WRITE_ONCE(*x, 1);
+}
+
+exists (x=2 /\ 1:r0=1)
diff --git a/tools/memory-model/litmus-tests/SB+fencembonceonces.litmus b/tools/memory-model/litmus-tests/SB+fencembonceonces.litmus
new file mode 100644
index 0000000000..ed5fff18d2
--- /dev/null
+++ b/tools/memory-model/litmus-tests/SB+fencembonceonces.litmus
@@ -0,0 +1,32 @@
+C SB+fencembonceonces
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates that full memory barriers suffice to
+ * order the store-buffering pattern, where each process writes to the
+ * variable that the preceding process reads.  (Locking and RCU can also
+ * suffice, but not much else.)
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	int r0;
+
+	WRITE_ONCE(*x, 1);
+	smp_mb();
+	r0 = READ_ONCE(*y);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	WRITE_ONCE(*y, 1);
+	smp_mb();
+	r0 = READ_ONCE(*x);
+}
+
+exists (0:r0=0 /\ 1:r0=0)
diff --git a/tools/memory-model/litmus-tests/SB+poonceonces.litmus b/tools/memory-model/litmus-tests/SB+poonceonces.litmus
new file mode 100644
index 0000000000..10d550730b
--- /dev/null
+++ b/tools/memory-model/litmus-tests/SB+poonceonces.litmus
@@ -0,0 +1,29 @@
+C SB+poonceonces
+
+(*
+ * Result: Sometimes
+ *
+ * This litmus test demonstrates that at least some ordering is required
+ * to order the store-buffering pattern, where each process writes to the
+ * variable that the preceding process reads.
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	int r0;
+
+	WRITE_ONCE(*x, 1);
+	r0 = READ_ONCE(*y);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	WRITE_ONCE(*y, 1);
+	r0 = READ_ONCE(*x);
+}
+
+exists (0:r0=0 /\ 1:r0=0)
diff --git a/tools/memory-model/litmus-tests/SB+rfionceonce-poonceonces.litmus b/tools/memory-model/litmus-tests/SB+rfionceonce-poonceonces.litmus
new file mode 100644
index 0000000000..04a1660366
--- /dev/null
+++ b/tools/memory-model/litmus-tests/SB+rfionceonce-poonceonces.litmus
@@ -0,0 +1,32 @@
+C SB+rfionceonce-poonceonces
+
+(*
+ * Result: Sometimes
+ *
+ * This litmus test demonstrates that LKMM is not fully multicopy atomic.
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	int r1;
+	int r2;
+
+	WRITE_ONCE(*x, 1);
+	r1 = READ_ONCE(*x);
+	r2 = READ_ONCE(*y);
+}
+
+P1(int *x, int *y)
+{
+	int r3;
+	int r4;
+
+	WRITE_ONCE(*y, 1);
+	r3 = READ_ONCE(*y);
+	r4 = READ_ONCE(*x);
+}
+
+locations [0:r1; 1:r3; x; y] (* Debug aid: Print things not in "exists". *)
+exists (0:r2=0 /\ 1:r4=0)
diff --git a/tools/memory-model/litmus-tests/WRC+poonceonces+Once.litmus b/tools/memory-model/litmus-tests/WRC+poonceonces+Once.litmus
new file mode 100644
index 0000000000..6a2bc12a1a
--- /dev/null
+++ b/tools/memory-model/litmus-tests/WRC+poonceonces+Once.litmus
@@ -0,0 +1,35 @@
+C WRC+poonceonces+Once
+
+(*
+ * Result: Sometimes
+ *
+ * This litmus test is an extension of the message-passing pattern,
+ * where the first write is moved to a separate process.  Note that this
+ * test has no ordering at all.
+ *)
+
+{}
+
+P0(int *x)
+{
+	WRITE_ONCE(*x, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*x);
+	WRITE_ONCE(*y, 1);
+}
+
+P2(int *x, int *y)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*y);
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 2:r0=1 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/WRC+pooncerelease+fencermbonceonce+Once.litmus b/tools/memory-model/litmus-tests/WRC+pooncerelease+fencermbonceonce+Once.litmus
new file mode 100644
index 0000000000..e9947250d7
--- /dev/null
+++ b/tools/memory-model/litmus-tests/WRC+pooncerelease+fencermbonceonce+Once.litmus
@@ -0,0 +1,38 @@
+C WRC+pooncerelease+fencermbonceonce+Once
+
+(*
+ * Result: Never
+ *
+ * This litmus test is an extension of the message-passing pattern, where
+ * the first write is moved to a separate process.  Because it features
+ * a release and a read memory barrier, it should be forbidden.  More
+ * specifically, this litmus test is forbidden because smp_store_release()
+ * is A-cumulative in LKMM.
+ *)
+
+{}
+
+P0(int *x)
+{
+	WRITE_ONCE(*x, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r0;
+
+	r0 = READ_ONCE(*x);
+	smp_store_release(y, 1);
+}
+
+P2(int *x, int *y)
+{
+	int r0;
+	int r1;
+
+	r0 = READ_ONCE(*y);
+	smp_rmb();
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ 2:r0=1 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/Z6.0+pooncelock+poonceLock+pombonce.litmus b/tools/memory-model/litmus-tests/Z6.0+pooncelock+poonceLock+pombonce.litmus
new file mode 100644
index 0000000000..415248fb66
--- /dev/null
+++ b/tools/memory-model/litmus-tests/Z6.0+pooncelock+poonceLock+pombonce.litmus
@@ -0,0 +1,42 @@
+C Z6.0+pooncelock+poonceLock+pombonce
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates how smp_mb__after_spinlock() may be
+ * used to ensure that accesses in different critical sections for a
+ * given lock running on different CPUs are nevertheless seen in order
+ * by CPUs not holding that lock.
+ *)
+
+{}
+
+P0(int *x, int *y, spinlock_t *mylock)
+{
+	spin_lock(mylock);
+	WRITE_ONCE(*x, 1);
+	WRITE_ONCE(*y, 1);
+	spin_unlock(mylock);
+}
+
+P1(int *y, int *z, spinlock_t *mylock)
+{
+	int r0;
+
+	spin_lock(mylock);
+	smp_mb__after_spinlock();
+	r0 = READ_ONCE(*y);
+	WRITE_ONCE(*z, 1);
+	spin_unlock(mylock);
+}
+
+P2(int *x, int *z)
+{
+	int r1;
+
+	WRITE_ONCE(*z, 2);
+	smp_mb();
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ z=2 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/Z6.0+pooncelock+pooncelock+pombonce.litmus b/tools/memory-model/litmus-tests/Z6.0+pooncelock+pooncelock+pombonce.litmus
new file mode 100644
index 0000000000..10a2aa04cd
--- /dev/null
+++ b/tools/memory-model/litmus-tests/Z6.0+pooncelock+pooncelock+pombonce.litmus
@@ -0,0 +1,40 @@
+C Z6.0+pooncelock+pooncelock+pombonce
+
+(*
+ * Result: Sometimes
+ *
+ * This example demonstrates that a pair of accesses made by different
+ * processes each while holding a given lock will not necessarily be
+ * seen as ordered by a third process not holding that lock.
+ *)
+
+{}
+
+P0(int *x, int *y, spinlock_t *mylock)
+{
+	spin_lock(mylock);
+	WRITE_ONCE(*x, 1);
+	WRITE_ONCE(*y, 1);
+	spin_unlock(mylock);
+}
+
+P1(int *y, int *z, spinlock_t *mylock)
+{
+	int r0;
+
+	spin_lock(mylock);
+	r0 = READ_ONCE(*y);
+	WRITE_ONCE(*z, 1);
+	spin_unlock(mylock);
+}
+
+P2(int *x, int *z)
+{
+	int r1;
+
+	WRITE_ONCE(*z, 2);
+	smp_mb();
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ z=2 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/Z6.0+pooncerelease+poacquirerelease+fencembonceonce.litmus b/tools/memory-model/litmus-tests/Z6.0+pooncerelease+poacquirerelease+fencembonceonce.litmus
new file mode 100644
index 0000000000..88e70b87a6
--- /dev/null
+++ b/tools/memory-model/litmus-tests/Z6.0+pooncerelease+poacquirerelease+fencembonceonce.litmus
@@ -0,0 +1,42 @@
+C Z6.0+pooncerelease+poacquirerelease+fencembonceonce
+
+(*
+ * Result: Sometimes
+ *
+ * This litmus test shows that a release-acquire chain, while sufficient
+ * when there is but one non-reads-from (AKA non-rf) link, does not suffice
+ * if there is more than one.  Of the three processes, only P1() reads from
+ * P0's write, which means that there are two non-rf links: P1() to P2()
+ * is a write-to-write link (AKA a "coherence" or just "co" link) and P2()
+ * to P0() is a read-to-write link (AKA a "from-reads" or just "fr" link).
+ * When there are two or more non-rf links, you typically will need one
+ * full barrier for each non-rf link.  (Exceptions include some cases
+ * involving locking.)
+ *)
+
+{}
+
+P0(int *x, int *y)
+{
+	WRITE_ONCE(*x, 1);
+	smp_store_release(y, 1);
+}
+
+P1(int *y, int *z)
+{
+	int r0;
+
+	r0 = smp_load_acquire(y);
+	smp_store_release(z, 1);
+}
+
+P2(int *x, int *z)
+{
+	int r1;
+
+	WRITE_ONCE(*z, 2);
+	smp_mb();
+	r1 = READ_ONCE(*x);
+}
+
+exists (1:r0=1 /\ z=2 /\ 2:r1=0)
diff --git a/tools/memory-model/litmus-tests/dep+plain.litmus b/tools/memory-model/litmus-tests/dep+plain.litmus
new file mode 100644
index 0000000000..ebf84daa9a
--- /dev/null
+++ b/tools/memory-model/litmus-tests/dep+plain.litmus
@@ -0,0 +1,31 @@
+C dep+plain
+
+(*
+ * Result: Never
+ *
+ * This litmus test demonstrates that in LKMM, plain accesses
+ * carry dependencies much like accesses to registers:
+ * The data stored to *z1 and *z2 by P0() originates from P0()'s
+ * READ_ONCE(), and therefore using that data to compute the
+ * conditional of P0()'s if-statement creates a control dependency
+ * from that READ_ONCE() to P0()'s WRITE_ONCE().
+ *)
+
+{}
+
+P0(int *x, int *y, int *z1, int *z2)
+{
+	int a = READ_ONCE(*x);
+	*z1 = a;
+	*z2 = *z1;
+	if (*z2 == 1)
+		WRITE_ONCE(*y, 1);
+}
+
+P1(int *x, int *y)
+{
+	int r = smp_load_acquire(y);
+	smp_store_release(x, r);
+}
+
+exists (x=1 /\ y=1)
diff --git a/tools/memory-model/lock.cat b/tools/memory-model/lock.cat
new file mode 100644
index 0000000000..53b5a49273
--- /dev/null
+++ b/tools/memory-model/lock.cat
@@ -0,0 +1,143 @@
+// SPDX-License-Identifier: GPL-2.0+
+(*
+ * Copyright (C) 2016 Luc Maranget <luc.maranget@inria.fr> for Inria
+ * Copyright (C) 2017 Alan Stern <stern@rowland.harvard.edu>
+ *)
+
+(*
+ * Generate coherence orders and handle lock operations
+ *)
+
+include "cross.cat"
+
+(*
+ * The lock-related events generated by herd7 are as follows:
+ *
+ * LKR		Lock-Read: the read part of a spin_lock() or successful
+ *			spin_trylock() read-modify-write event pair
+ * LKW		Lock-Write: the write part of a spin_lock() or successful
+ *			spin_trylock() RMW event pair
+ * UL		Unlock: a spin_unlock() event
+ * LF		Lock-Fail: a failed spin_trylock() event
+ * RL		Read-Locked: a spin_is_locked() event which returns True
+ * RU		Read-Unlocked: a spin_is_locked() event which returns False
+ *
+ * LKR and LKW events always come paired, like all RMW event sequences.
+ *
+ * LKR, LF, RL, and RU are read events; LKR has Acquire ordering.
+ * LKW and UL are write events; UL has Release ordering.
+ * LKW, LF, RL, and RU have no ordering properties.
+ *)
+
+(* Backward compatibility *)
+let RL = try RL with emptyset
+let RU = try RU with emptyset
+
+(* Treat RL as a kind of LF: a read with no ordering properties *)
+let LF = LF | RL
+
+(* There should be no ordinary R or W accesses to spinlocks or SRCU structs *)
+let ALL-LOCKS = LKR | LKW | UL | LF | RU | Srcu-lock | Srcu-unlock | Sync-srcu
+flag ~empty [M \ IW \ ALL-LOCKS] ; loc ; [ALL-LOCKS] as mixed-lock-accesses
+
+(* Link Lock-Reads to their RMW-partner Lock-Writes *)
+let lk-rmw = ([LKR] ; po-loc ; [LKW]) \ (po ; po)
+let rmw = rmw | lk-rmw
+
+(* The litmus test is invalid if an LKR/LKW event is not part of an RMW pair *)
+flag ~empty LKW \ range(lk-rmw) as unpaired-LKW
+flag ~empty LKR \ domain(lk-rmw) as unpaired-LKR
+
+(*
+ * An LKR must always see an unlocked value; spin_lock() calls nested
+ * inside a critical section (for the same lock) always deadlock.
+ *)
+empty ([LKW] ; po-loc ; [LKR]) \ (po-loc ; [UL] ; po-loc) as lock-nest
+
+(* The final value of a spinlock should not be tested *)
+flag ~empty [FW] ; loc ; [ALL-LOCKS] as lock-final
+
+(*
+ * Put lock operations in their appropriate classes, but leave UL out of W
+ * until after the co relation has been generated.
+ *)
+let R = R | LKR | LF | RU
+let W = W | LKW
+
+let Release = Release | UL
+let Acquire = Acquire | LKR
+
+(* Match LKW events to their corresponding UL events *)
+let critical = ([LKW] ; po-loc ; [UL]) \ (po-loc ; [LKW | UL] ; po-loc)
+
+flag ~empty UL \ range(critical) as unmatched-unlock
+
+(* Allow up to one unmatched LKW per location; more must deadlock *)
+let UNMATCHED-LKW = LKW \ domain(critical)
+empty ([UNMATCHED-LKW] ; loc ; [UNMATCHED-LKW]) \ id as unmatched-locks
+
+(* rfi for LF events: link each LKW to the LF events in its critical section *)
+let rfi-lf = ([LKW] ; po-loc ; [LF]) \ ([LKW] ; po-loc ; [UL] ; po-loc)
+
+(* rfe for LF events *)
+let all-possible-rfe-lf =
+	(*
+	 * Given an LF event r, compute the possible rfe edges for that event
+	 * (all those starting from LKW events in other threads),
+	 * and then convert that relation to a set of single-edge relations.
+	 *)
+	let possible-rfe-lf r =
+		let pair-to-relation p = p ++ 0
+		in map pair-to-relation ((LKW * {r}) & loc & ext)
+	(* Do this for each LF event r that isn't in rfi-lf *)
+	in map possible-rfe-lf (LF \ range(rfi-lf))
+
+(* Generate all rf relations for LF events *)
+with rfe-lf from cross(all-possible-rfe-lf)
+let rf-lf = rfe-lf | rfi-lf
+
+(*
+ * RU, i.e., spin_is_locked() returning False, is slightly different.
+ * We rely on the memory model to rule out cases where spin_is_locked()
+ * within one of the lock's critical sections returns False.
+ *)
+
+(* rfi for RU events: an RU may read from the last po-previous UL *)
+let rfi-ru = ([UL] ; po-loc ; [RU]) \ ([UL] ; po-loc ; [LKW] ; po-loc)
+
+(* rfe for RU events: an RU may read from an external UL or the initial write *)
+let all-possible-rfe-ru =
+	let possible-rfe-ru r =
+		let pair-to-relation p = p ++ 0
+		in map pair-to-relation (((UL | IW) * {r}) & loc & ext)
+	in map possible-rfe-ru RU
+
+(* Generate all rf relations for RU events *)
+with rfe-ru from cross(all-possible-rfe-ru)
+let rf-ru = rfe-ru | rfi-ru
+
+(* Final rf relation *)
+let rf = rf | rf-lf | rf-ru
+
+(* Generate all co relations, including LKW events but not UL *)
+let co0 = co0 | ([IW] ; loc ; [LKW]) |
+	(([LKW] ; loc ; [UNMATCHED-LKW]) \ [UNMATCHED-LKW])
+include "cos-opt.cat"
+let W = W | UL
+let M = R | W
+
+(* Merge UL events into co *)
+let co = (co | critical | (critical^-1 ; co))+
+let coe = co & ext
+let coi = co & int
+
+(* Merge LKR events into rf *)
+let rf = rf | ([IW | UL] ; singlestep(co) ; lk-rmw^-1)
+let rfe = rf & ext
+let rfi = rf & int
+
+let fr = rf^-1 ; co
+let fre = fr & ext
+let fri = fr & int
+
+show co,rf,fr
diff --git a/tools/memory-model/scripts/README b/tools/memory-model/scripts/README
new file mode 100644
index 0000000000..fb39bd0fd1
--- /dev/null
+++ b/tools/memory-model/scripts/README
@@ -0,0 +1,110 @@
+			============
+			LKMM SCRIPTS
+			============
+
+
+These scripts are run from the tools/memory-model directory.
+
+checkalllitmus.sh
+
+	Run all litmus tests in the litmus-tests directory, checking
+	the results against the expected results recorded in the
+	"Result:" comment lines.
+
+checkghlitmus.sh
+
+	Run all litmus tests in the https://github.com/paulmckrcu/litmus
+	archive that are C-language and that have "Result:" comment lines
+	documenting expected results, comparing the actual results to
+	those expected.
+
+checklitmushist.sh
+
+	Run all litmus tests having .litmus.out files from previous
+	initlitmushist.sh or newlitmushist.sh runs, comparing the
+	herd7 output to that of the original runs.
+
+checklitmus.sh
+
+	Check a single litmus test against its "Result:" expected result.
+	Not intended to for manual use.
+
+checktheselitmus.sh
+
+	Check the specified list of litmus tests against their "Result:"
+	expected results.  This takes optional parseargs.sh arguments,
+	followed by "--" followed by pathnames starting from the current
+	directory.
+
+cmplitmushist.sh
+
+	Compare output from two different runs of the same litmus tests,
+	with the absolute pathnames of the tests to run provided one
+	name per line on standard input.  Not normally run manually,
+	provided instead for use by other scripts.
+
+initlitmushist.sh
+
+	Run all litmus tests having no more than the specified number
+	of processes given a specified timeout, recording the results
+	in .litmus.out files.
+
+judgelitmus.sh
+
+	Given a .litmus file and its herd7 output, check the output file
+	against the .litmus file's "Result:" comment to judge whether
+	the test ran correctly.  Not normally run manually, provided
+	instead for use by other scripts.
+
+newlitmushist.sh
+
+	For all new or updated litmus tests having no more than the
+	specified number of processes given a specified timeout, run
+	and record the results in .litmus.out files.
+
+parseargs.sh
+
+	Parse command-line arguments.  Not normally run manually,
+	provided instead for use by other scripts.
+
+runlitmushist.sh
+
+	Run the litmus tests whose absolute pathnames are provided one
+	name per line on standard input.  Not normally run manually,
+	provided instead for use by other scripts.
+
+README
+
+	This file
+
+Testing a change to LKMM might go as follows:
+
+	# Populate expected results without that change, and
+	# runs for about an hour on an 8-CPU x86 system:
+	scripts/initlitmushist.sh --timeout 10m --procs 10
+	# Incorporate the change:
+	git am -s -3 /path/to/patch # Or whatever it takes.
+
+	# Test the new version of LKMM as follows...
+
+	# Runs in seconds, good smoke test:
+	scripts/checkalllitmus.sh
+
+	# Compares results to those produced by initlitmushist.sh,
+	# and runs for about an hour on an 8-CPU x86 system:
+	scripts/checklitmushist.sh --timeout 10m --procs 10
+
+	# Checks results against Result tags, runs in minutes:
+	scripts/checkghlitmus.sh --timeout 10m --procs 10
+
+The checkghlitmus.sh should not report errors in cases where the
+checklitmushist.sh script did not also report a change.  However,
+this check is nevertheless valuable because it can find errors in the
+original version of LKMM.  Note however, that given the above procedure,
+an error in the original LKMM version that is fixed by the patch will
+be reported both as a mismatch by checklitmushist.sh and as an error
+by checkghlitmus.sh.  One exception to this rule of thumb is when the
+test fails completely on the original version of LKMM and passes on the
+new version.  In this case, checklitmushist.sh will report a mismatch
+and checkghlitmus.sh will report success.  This happens when the change
+to LKMM introduces a new primitive for which litmus tests already existed.
diff --git a/tools/memory-model/scripts/checkalllitmus.sh b/tools/memory-model/scripts/checkalllitmus.sh
new file mode 100755
index 0000000000..2d3ee850a8
--- /dev/null
+++ b/tools/memory-model/scripts/checkalllitmus.sh
@@ -0,0 +1,65 @@
+#!/bin/bash
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Run herd7 tests on all .litmus files in the litmus-tests directory
+# and check each file's result against a "Result:" comment within that
+# litmus test.  If the verification result does not match that specified
+# in the litmus test, this script prints an error message prefixed with
+# "^^^".  It also outputs verification results to a file whose name is
+# that of the specified litmus test, but with ".out" appended.
+#
+# If the --hw argument is specified, this script translates the .litmus
+# C-language file to the specified type of assembly and verifies that.
+# But in this case, litmus tests using complex synchronization (such as
+# locking, RCU, and SRCU) are cheerfully ignored.
+#
+# Usage:
+#	checkalllitmus.sh
+#
+# Run this in the directory containing the memory model.
+#
+# This script makes no attempt to run the litmus tests concurrently.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+. scripts/parseargs.sh
+
+litmusdir=litmus-tests
+if test -d "$litmusdir" -a -r "$litmusdir" -a -x "$litmusdir"
+then
+	:
+else
+	echo ' --- ' error: $litmusdir is not an accessible directory
+	exit 255
+fi
+
+# Create any new directories that have appeared in the litmus-tests
+# directory since the last run.
+if test "$LKMM_DESTDIR" != "."
+then
+	find $litmusdir -type d -print |
+	( cd "$LKMM_DESTDIR"; sed -e 's/^/mkdir -p /' | sh )
+fi
+
+# Run the script on all the litmus tests in the specified directory
+ret=0
+for i in $litmusdir/*.litmus
+do
+	if test -n "$LKMM_HW_MAP_FILE" && ! scripts/simpletest.sh $i
+	then
+		continue
+	fi
+	if ! scripts/checklitmus.sh $i
+	then
+		ret=1
+	fi
+done
+if test "$ret" -ne 0
+then
+	echo " ^^^ VERIFICATION MISMATCHES" 1>&2
+else
+	echo All litmus tests verified as was expected. 1>&2
+fi
+exit $ret
diff --git a/tools/memory-model/scripts/checkghlitmus.sh b/tools/memory-model/scripts/checkghlitmus.sh
new file mode 100755
index 0000000000..d3dfb32125
--- /dev/null
+++ b/tools/memory-model/scripts/checkghlitmus.sh
@@ -0,0 +1,66 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Runs the C-language litmus tests having a maximum number of processes
+# to run, defaults to 6.
+#
+# sh checkghlitmus.sh
+#
+# Run from the Linux kernel tools/memory-model directory.  See the
+# parseargs.sh scripts for arguments.
+
+. scripts/parseargs.sh
+. scripts/hwfnseg.sh
+
+T=/tmp/checkghlitmus.sh.$$
+trap 'rm -rf $T' 0
+mkdir $T
+
+# Clone the repository if it is not already present.
+if test -d litmus
+then
+	:
+else
+	git clone https://github.com/paulmckrcu/litmus
+	( cd litmus; git checkout origin/master )
+fi
+
+# Create any new directories that have appeared in the github litmus
+# repo since the last run.
+if test "$LKMM_DESTDIR" != "."
+then
+	find litmus -type d -print |
+	( cd "$LKMM_DESTDIR"; sed -e 's/^/mkdir -p /' | sh )
+fi
+
+# Create a list of the specified litmus tests previously run.
+( cd $LKMM_DESTDIR; find litmus -name "*.litmus${hwfnseg}.out" -print ) |
+	sed -e "s/${hwfnseg}"'\.out$//' |
+	xargs -r grep -E -l '^ \* Result: (Never|Sometimes|Always|DEADLOCK)' |
+	xargs -r grep -L "^P${LKMM_PROCS}"> $T/list-C-already
+
+# Create a list of C-language litmus tests with "Result:" commands and
+# no more than the specified number of processes.
+find litmus -name '*.litmus' -print | mselect7 -arch C > $T/list-C
+xargs < $T/list-C -r grep -E -l '^ \* Result: (Never|Sometimes|Always|DEADLOCK)' > $T/list-C-result
+xargs < $T/list-C-result -r grep -L "^P${LKMM_PROCS}" > $T/list-C-result-short
+
+# Form list of tests without corresponding .out files
+sort $T/list-C-already $T/list-C-result-short | uniq -u > $T/list-C-needed
+
+# Run any needed tests.
+if scripts/runlitmushist.sh < $T/list-C-needed > $T/run.stdout 2> $T/run.stderr
+then
+	errs=
+else
+	errs=1
+fi
+
+sed < $T/list-C-result-short -e 's,^,scripts/judgelitmus.sh ,' |
+	sh > $T/judge.stdout 2> $T/judge.stderr
+
+if test -n "$errs"
+then
+	cat $T/run.stderr 1>&2
+fi
+grep '!!!' $T/judge.stdout
diff --git a/tools/memory-model/scripts/checklitmus.sh b/tools/memory-model/scripts/checklitmus.sh
new file mode 100755
index 0000000000..4c1d0cf0dd
--- /dev/null
+++ b/tools/memory-model/scripts/checklitmus.sh
@@ -0,0 +1,19 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Invokes runlitmus.sh and judgelitmus.sh on its arguments to run the
+# specified litmus test and pass judgment on the results.
+#
+# Usage:
+#	checklitmus.sh file.litmus
+#
+# Run this in the directory containing the memory model, specifying the
+# pathname of the litmus test to check.  The caller is expected to have
+# properly set up the LKMM environment variables.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+scripts/runlitmus.sh $1
+scripts/judgelitmus.sh $1
diff --git a/tools/memory-model/scripts/checklitmushist.sh b/tools/memory-model/scripts/checklitmushist.sh
new file mode 100755
index 0000000000..406ecfc0ae
--- /dev/null
+++ b/tools/memory-model/scripts/checklitmushist.sh
@@ -0,0 +1,60 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Reruns the C-language litmus tests previously run that match the
+# specified criteria, and compares the result to that of the previous
+# runs from initlitmushist.sh and/or newlitmushist.sh.
+#
+# sh checklitmushist.sh
+#
+# Run from the Linux kernel tools/memory-model directory.
+# See scripts/parseargs.sh for list of arguments.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+. scripts/parseargs.sh
+
+T=/tmp/checklitmushist.sh.$$
+trap 'rm -rf $T' 0
+mkdir $T
+
+if test -d litmus
+then
+	:
+else
+	echo Run scripts/initlitmushist.sh first, need litmus repo.
+	exit 1
+fi
+
+# Create the results directory and populate it with subdirectories.
+# The initial output is created here to avoid clobbering the output
+# generated earlier.
+mkdir $T/results
+find litmus -type d -print | ( cd $T/results; sed -e 's/^/mkdir -p /' | sh )
+
+# Create the list of litmus tests already run, then remove those that
+# are excluded by this run's --procs argument.
+( cd $LKMM_DESTDIR; find litmus -name '*.litmus.out' -print ) |
+	sed -e 's/\.out$//' |
+	xargs -r grep -L "^P${LKMM_PROCS}"> $T/list-C-already
+xargs < $T/list-C-already -r grep -L "^P${LKMM_PROCS}" > $T/list-C-short
+
+# Redirect output, run tests, then restore destination directory.
+destdir="$LKMM_DESTDIR"
+LKMM_DESTDIR=$T/results; export LKMM_DESTDIR
+scripts/runlitmushist.sh < $T/list-C-short > $T/runlitmushist.sh.out 2>&1
+LKMM_DESTDIR="$destdir"; export LKMM_DESTDIR
+
+# Move the newly generated .litmus.out files to .litmus.out.new files
+# in the destination directory.
+cdir=`pwd`
+ddir=`awk -v c="$cdir" -v d="$LKMM_DESTDIR" \
+	'END { if (d ~ /^\//) print d; else print c "/" d; }' < /dev/null`
+( cd $T/results; find litmus -type f -name '*.litmus.out' -print |
+  sed -e 's,^.*$,cp & '"$ddir"'/&.new,' | sh )
+
+sed < $T/list-C-short -e 's,^,'"$LKMM_DESTDIR/"',' |
+	sh scripts/cmplitmushist.sh
+exit $?
diff --git a/tools/memory-model/scripts/checktheselitmus.sh b/tools/memory-model/scripts/checktheselitmus.sh
new file mode 100755
index 0000000000..10eeb5ecea
--- /dev/null
+++ b/tools/memory-model/scripts/checktheselitmus.sh
@@ -0,0 +1,43 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Invokes checklitmus.sh on its arguments to run the specified litmus
+# test and pass judgment on the results.
+#
+# Usage:
+#	checktheselitmus.sh -- [ file1.litmus [ file2.litmus ... ] ]
+#
+# Run this in the directory containing the memory model, specifying the
+# pathname of the litmus test to check.  The usual parseargs.sh arguments
+# can be specified prior to the "--".
+#
+# This script is intended for use with pathnames that start from the
+# tools/memory-model directory.  If some of the pathnames instead start at
+# the root directory, they all must do so and the "--destdir /" parseargs.sh
+# argument must be specified prior to the "--".  Alternatively, some other
+# "--destdir" argument can be supplied as long as the needed subdirectories
+# are populated.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+. scripts/parseargs.sh
+
+ret=0
+for i in "$@"
+do
+	if scripts/checklitmus.sh $i
+	then
+		:
+	else
+		ret=1
+	fi
+done
+if test "$ret" -ne 0
+then
+	echo " ^^^ VERIFICATION MISMATCHES" 1>&2
+else
+	echo All litmus tests verified as was expected. 1>&2
+fi
+exit $ret
diff --git a/tools/memory-model/scripts/cmplitmushist.sh b/tools/memory-model/scripts/cmplitmushist.sh
new file mode 100755
index 0000000000..ca1ac8b646
--- /dev/null
+++ b/tools/memory-model/scripts/cmplitmushist.sh
@@ -0,0 +1,132 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Compares .out and .out.new files for each name on standard input,
+# one full pathname per line.  Outputs comparison results followed by
+# a summary.
+#
+# sh cmplitmushist.sh
+
+T=/tmp/cmplitmushist.sh.$$
+trap 'rm -rf $T' 0
+mkdir $T
+
+# comparetest oldpath newpath
+badmacnam=0
+timedout=0
+perfect=0
+obsline=0
+noobsline=0
+obsresult=0
+badcompare=0
+comparetest () {
+	if grep -q ': Unknown macro ' $1 || grep -q ': Unknown macro ' $2
+	then
+		if grep -q ': Unknown macro ' $1
+		then
+			badname=`grep ': Unknown macro ' $1 |
+				sed -e 's/^.*: Unknown macro //' |
+				sed -e 's/ (User error).*$//'`
+			echo 'Current LKMM version does not know "'$badname'"' $1
+		fi
+		if grep -q ': Unknown macro ' $2
+		then
+			badname=`grep ': Unknown macro ' $2 |
+				sed -e 's/^.*: Unknown macro //' |
+				sed -e 's/ (User error).*$//'`
+			echo 'Current LKMM version does not know "'$badname'"' $2
+		fi
+		badmacnam=`expr "$badmacnam" + 1`
+		return 0
+	elif grep -q '^Command exited with non-zero status 124' $1 ||
+	     grep -q '^Command exited with non-zero status 124' $2
+	then
+		if grep -q '^Command exited with non-zero status 124' $1 &&
+		   grep -q '^Command exited with non-zero status 124' $2
+		then
+			echo Both runs timed out: $2
+		elif grep -q '^Command exited with non-zero status 124' $1
+		then
+			echo Old run timed out: $2
+		elif grep -q '^Command exited with non-zero status 124' $2
+		then
+			echo New run timed out: $2
+		fi
+		timedout=`expr "$timedout" + 1`
+		return 0
+	fi
+	grep -v 'maxresident)k\|minor)pagefaults\|^Time' $1 > $T/oldout
+	grep -v 'maxresident)k\|minor)pagefaults\|^Time' $2 > $T/newout
+	if cmp -s $T/oldout $T/newout && grep -q '^Observation' $1
+	then
+		echo Exact output match: $2
+		perfect=`expr "$perfect" + 1`
+		return 0
+	fi
+
+	grep '^Observation' $1 > $T/oldout
+	grep '^Observation' $2 > $T/newout
+	if test -s $T/oldout -o -s $T/newout
+	then
+		if cmp -s $T/oldout $T/newout
+		then
+			echo Matching Observation result and counts: $2
+			obsline=`expr "$obsline" + 1`
+			return 0
+		fi
+	else
+		echo Missing Observation line "(e.g., syntax error)": $2
+		noobsline=`expr "$noobsline" + 1`
+		return 0
+	fi
+
+	grep '^Observation' $1 | awk '{ print $3 }' > $T/oldout
+	grep '^Observation' $2 | awk '{ print $3 }' > $T/newout
+	if cmp -s $T/oldout $T/newout
+	then
+		echo Matching Observation Always/Sometimes/Never result: $2
+		obsresult=`expr "$obsresult" + 1`
+		return 0
+	fi
+	echo ' !!!' Result changed: $2
+	badcompare=`expr "$badcompare" + 1`
+	return 1
+}
+
+sed -e 's/^.*$/comparetest &.out &.out.new/' > $T/cmpscript
+. $T/cmpscript > $T/cmpscript.out
+cat $T/cmpscript.out
+
+echo ' ---' Summary: 1>&2
+grep '!!!' $T/cmpscript.out 1>&2
+if test "$perfect" -ne 0
+then
+	echo Exact output matches: $perfect 1>&2
+fi
+if test "$obsline" -ne 0
+then
+	echo Matching Observation result and counts: $obsline 1>&2
+fi
+if test "$noobsline" -ne 0
+then
+	echo Missing Observation line "(e.g., syntax error)": $noobsline 1>&2
+fi
+if test "$obsresult" -ne 0
+then
+	echo Matching Observation Always/Sometimes/Never result: $obsresult 1>&2
+fi
+if test "$timedout" -ne 0
+then
+	echo "!!!" Timed out: $timedout 1>&2
+fi
+if test "$badmacnam" -ne 0
+then
+	echo "!!!" Unknown primitive: $badmacnam 1>&2
+fi
+if test "$badcompare" -ne 0
+then
+	echo "!!!" Result changed: $badcompare 1>&2
+	exit 1
+fi
+
+exit 0
diff --git a/tools/memory-model/scripts/hwfnseg.sh b/tools/memory-model/scripts/hwfnseg.sh
new file mode 100755
index 0000000000..580c328118
--- /dev/null
+++ b/tools/memory-model/scripts/hwfnseg.sh
@@ -0,0 +1,20 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Generate the hardware extension to the litmus-test filename, or the
+# empty string if this is an LKMM run.  The extension is placed in
+# the shell variable hwfnseg.
+#
+# Usage:
+#	. hwfnseg.sh
+#
+# Copyright IBM Corporation, 2019
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+if test -z "$LKMM_HW_MAP_FILE"
+then
+	hwfnseg=
+else
+	hwfnseg=".$LKMM_HW_MAP_FILE"
+fi
diff --git a/tools/memory-model/scripts/initlitmushist.sh b/tools/memory-model/scripts/initlitmushist.sh
new file mode 100755
index 0000000000..31ea782955
--- /dev/null
+++ b/tools/memory-model/scripts/initlitmushist.sh
@@ -0,0 +1,68 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Runs the C-language litmus tests matching the specified criteria.
+# Generates the output for each .litmus file into a corresponding
+# .litmus.out file, and does not judge the result.
+#
+# sh initlitmushist.sh
+#
+# Run from the Linux kernel tools/memory-model directory.
+# See scripts/parseargs.sh for list of arguments.
+#
+# This script can consume significant wallclock time and CPU, especially as
+# the value of --procs rises.  On a four-core (eight hardware threads)
+# 2.5GHz x86 with a one-minute per-run timeout:
+#
+# --procs wallclock CPU		timeouts	tests
+#	1 0m11.241s 0m1.086s           0	   19
+#	2 1m12.598s 2m8.459s           2	  393
+#	3 1m30.007s 6m2.479s           4	 2291
+#	4 3m26.042s 18m5.139s	       9	 3217
+#	5 4m26.661s 23m54.128s	      13	 3784
+#	6 4m41.900s 26m4.721s         13	 4352
+#	7 5m51.463s 35m50.868s        13	 4626
+#	8 10m5.235s 68m43.672s        34	 5117
+#	9 15m57.80s 105m58.101s       69	 5156
+#      10 16m14.13s 103m35.009s       69         5165
+#      20 27m48.55s 198m3.286s       156         5269
+#
+# Increasing the timeout on the 20-process run to five minutes increases
+# the runtime to about 90 minutes with the CPU time rising to about
+# 10 hours.  On the other hand, it decreases the number of timeouts to 101.
+#
+# Note that there are historical tests for which herd7 will fail
+# completely, for example, litmus/manual/atomic/C-unlock-wait-00.litmus
+# contains a call to spin_unlock_wait(), which no longer exists in either
+# the kernel or LKMM.
+
+. scripts/parseargs.sh
+
+T=/tmp/initlitmushist.sh.$$
+trap 'rm -rf $T' 0
+mkdir $T
+
+if test -d litmus
+then
+	:
+else
+	git clone https://github.com/paulmckrcu/litmus
+	( cd litmus; git checkout origin/master )
+fi
+
+# Create any new directories that have appeared in the github litmus
+# repo since the last run.
+if test "$LKMM_DESTDIR" != "."
+then
+	find litmus -type d -print |
+	( cd "$LKMM_DESTDIR"; sed -e 's/^/mkdir -p /' | sh )
+fi
+
+# Create a list of the C-language litmus tests with no more than the
+# specified number of processes (per the --procs argument).
+find litmus -name '*.litmus' -print | mselect7 -arch C > $T/list-C
+xargs < $T/list-C -r grep -L "^P${LKMM_PROCS}" > $T/list-C-short
+
+scripts/runlitmushist.sh < $T/list-C-short
+
+exit 0
diff --git a/tools/memory-model/scripts/judgelitmus.sh b/tools/memory-model/scripts/judgelitmus.sh
new file mode 100755
index 0000000000..1ec5d89fcf
--- /dev/null
+++ b/tools/memory-model/scripts/judgelitmus.sh
@@ -0,0 +1,156 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Given a .litmus test and the corresponding litmus output file, check
+# the .litmus.out file against the "Result:" comment to judge whether the
+# test ran correctly.  If the --hw argument is omitted, check against the
+# LKMM output, which is assumed to be in file.litmus.out. If either a
+# "DATARACE" marker in the "Result:" comment or a "Flag data-race" marker
+# in the LKMM output is present, the other must also be as well, at least
+# for litmus tests having a "Result:" comment. In this case, a failure of
+# the Always/Sometimes/Never portion of the "Result:" prediction will be
+# noted, but forgiven.
+#
+# If the --hw argument is provided, this is assumed to be a hardware
+# test, and the output is assumed to be in file.litmus.HW.out, where
+# "HW" is the --hw argument. In addition, non-Sometimes verification
+# results will be noted, but forgiven.  Furthermore, if there is no
+# "Result:" comment but there is an LKMM .litmus.out file, the observation
+# in that file will be used to judge the assembly-language verification.
+#
+# Usage:
+#	judgelitmus.sh file.litmus
+#
+# Run this in the directory containing the memory model, specifying the
+# pathname of the litmus test to check.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+litmus=$1
+
+if test -f "$litmus" -a -r "$litmus"
+then
+	:
+else
+	echo ' --- ' error: \"$litmus\" is not a readable file
+	exit 255
+fi
+if test -z "$LKMM_HW_MAP_FILE"
+then
+	litmusout=$litmus.out
+	lkmmout=
+else
+	litmusout="`echo $litmus |
+		sed -e 's/\.litmus$/.litmus.'${LKMM_HW_MAP_FILE}'/'`.out"
+	lkmmout=$litmus.out
+fi
+if test -f "$LKMM_DESTDIR/$litmusout" -a -r "$LKMM_DESTDIR/$litmusout"
+then
+	:
+else
+	echo ' --- ' error: \"$LKMM_DESTDIR/$litmusout is not a readable file
+	exit 255
+fi
+if grep -q '^Flag data-race$' "$LKMM_DESTDIR/$litmusout"
+then
+	datarace_modeled=1
+fi
+if grep -q '^[( ]\* Result: ' $litmus
+then
+	outcome=`grep -m 1 '^[( ]\* Result: ' $litmus | awk '{ print $3 }'`
+	if grep -m1 '^[( ]\* Result: .* DATARACE' $litmus
+	then
+		datarace_predicted=1
+	fi
+	if test -n "$datarace_predicted" -a -z "$datarace_modeled" -a -z "$LKMM_HW_MAP_FILE"
+	then
+		echo '!!! Predicted data race not modeled' $litmus
+		exit 252
+	elif test -z "$datarace_predicted" -a -n "$datarace_modeled"
+	then
+		# Note that hardware models currently don't model data races
+		echo '!!! Unexpected data race modeled' $litmus
+		exit 253
+	fi
+elif test -n "$LKMM_HW_MAP_FILE" && grep -q '^Observation' $LKMM_DESTDIR/$lkmmout > /dev/null 2>&1
+then
+	outcome=`grep -m 1 '^Observation ' $LKMM_DESTDIR/$lkmmout | awk '{ print $3 }'`
+else
+	outcome=specified
+fi
+
+grep '^Observation' $LKMM_DESTDIR/$litmusout
+if grep -q '^Observation' $LKMM_DESTDIR/$litmusout
+then
+	:
+elif grep ': Unknown macro ' $LKMM_DESTDIR/$litmusout
+then
+	badname=`grep ': Unknown macro ' $LKMM_DESTDIR/$litmusout |
+		sed -e 's/^.*: Unknown macro //' |
+		sed -e 's/ (User error).*$//'`
+	badmsg=' !!! Current LKMM version does not know "'$badname'"'" $litmus"
+	echo $badmsg
+	if ! grep -q '!!!' $LKMM_DESTDIR/$litmusout
+	then
+		echo ' !!! '$badmsg >> $LKMM_DESTDIR/$litmusout 2>&1
+	fi
+	exit 254
+elif grep '^Command exited with non-zero status 124' $LKMM_DESTDIR/$litmusout
+then
+	echo ' !!! Timeout' $litmus
+	if ! grep -q '!!!' $LKMM_DESTDIR/$litmusout
+	then
+		echo ' !!! Timeout' >> $LKMM_DESTDIR/$litmusout 2>&1
+	fi
+	exit 124
+else
+	echo ' !!! Verification error' $litmus
+	if ! grep -q '!!!' $LKMM_DESTDIR/$litmusout
+	then
+		echo ' !!! Verification error' >> $LKMM_DESTDIR/$litmusout 2>&1
+	fi
+	exit 255
+fi
+if test "$outcome" = DEADLOCK
+then
+	if grep '^Observation' $LKMM_DESTDIR/$litmusout | grep -q 'Never 0 0$'
+	then
+		ret=0
+	else
+		echo " !!! Unexpected non-$outcome verification" $litmus
+		if ! grep -q '!!!' $LKMM_DESTDIR/$litmusout
+		then
+			echo " !!! Unexpected non-$outcome verification" >> $LKMM_DESTDIR/$litmusout 2>&1
+		fi
+		ret=1
+	fi
+elif grep '^Observation' $LKMM_DESTDIR/$litmusout | grep -q 'Never 0 0$'
+then
+	echo " !!! Unexpected non-$outcome deadlock" $litmus
+	if ! grep -q '!!!' $LKMM_DESTDIR/$litmusout
+	then
+		echo " !!! Unexpected non-$outcome deadlock" $litmus >> $LKMM_DESTDIR/$litmusout 2>&1
+	fi
+	ret=1
+elif grep '^Observation' $LKMM_DESTDIR/$litmusout | grep -q $outcome || test "$outcome" = Maybe
+then
+	ret=0
+else
+	if test \( -n "$LKMM_HW_MAP_FILE" -a "$outcome" = Sometimes \) -o -n "$datarace_modeled"
+	then
+		flag="--- Forgiven"
+		ret=0
+	else
+		flag="!!! Unexpected"
+		ret=1
+	fi
+	echo " $flag non-$outcome verification" $litmus
+	if ! grep -qe "$flag" $LKMM_DESTDIR/$litmusout
+	then
+		echo " $flag non-$outcome verification" >> $LKMM_DESTDIR/$litmusout 2>&1
+	fi
+fi
+tail -2 $LKMM_DESTDIR/$litmusout | head -1
+exit $ret
diff --git a/tools/memory-model/scripts/newlitmushist.sh b/tools/memory-model/scripts/newlitmushist.sh
new file mode 100755
index 0000000000..25235e2049
--- /dev/null
+++ b/tools/memory-model/scripts/newlitmushist.sh
@@ -0,0 +1,61 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Runs the C-language litmus tests matching the specified criteria
+# that do not already have a corresponding .litmus.out file, and does
+# not judge the result.
+#
+# sh newlitmushist.sh
+#
+# Run from the Linux kernel tools/memory-model directory.
+# See scripts/parseargs.sh for list of arguments.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+. scripts/parseargs.sh
+
+T=/tmp/newlitmushist.sh.$$
+trap 'rm -rf $T' 0
+mkdir $T
+
+if test -d litmus
+then
+	:
+else
+	echo Run scripts/initlitmushist.sh first, need litmus repo.
+	exit 1
+fi
+
+# Create any new directories that have appeared in the github litmus
+# repo since the last run.
+if test "$LKMM_DESTDIR" != "."
+then
+	find litmus -type d -print |
+	( cd "$LKMM_DESTDIR"; sed -e 's/^/mkdir -p /' | sh )
+fi
+
+# Create a list of the C-language litmus tests previously run.
+( cd $LKMM_DESTDIR; find litmus -name '*.litmus.out' -print ) |
+	sed -e 's/\.out$//' |
+	xargs -r grep -L "^P${LKMM_PROCS}"> $T/list-C-already
+
+# Form full list of litmus tests with no more than the specified
+# number of processes (per the --procs argument).
+find litmus -name '*.litmus' -print | mselect7 -arch C > $T/list-C-all
+xargs < $T/list-C-all -r grep -L "^P${LKMM_PROCS}" > $T/list-C-short
+
+# Form list of new tests.  Note: This does not handle litmus-test deletion!
+sort $T/list-C-already $T/list-C-short | uniq -u > $T/list-C-new
+
+# Form list of litmus tests that have changed since the last run.
+sed < $T/list-C-short -e 's,^.*$,if test & -nt '"$LKMM_DESTDIR"'/&.out; then echo &; fi,' > $T/list-C-script
+sh $T/list-C-script > $T/list-C-newer
+
+# Merge the list of new and of updated litmus tests: These must be (re)run.
+sort -u $T/list-C-new $T/list-C-newer > $T/list-C-needed
+
+scripts/runlitmushist.sh < $T/list-C-needed
+
+exit 0
diff --git a/tools/memory-model/scripts/parseargs.sh b/tools/memory-model/scripts/parseargs.sh
new file mode 100755
index 0000000000..08ded59098
--- /dev/null
+++ b/tools/memory-model/scripts/parseargs.sh
@@ -0,0 +1,147 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Parse arguments common to the various scripts.
+#
+# . scripts/parseargs.sh
+#
+# Include into other Linux kernel tools/memory-model scripts.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+T=/tmp/parseargs.sh.$$
+mkdir $T
+
+# Initialize one parameter: initparam name default
+initparam () {
+	echo if test -z '"$'$1'"' > $T/s
+	echo then >> $T/s
+	echo	$1='"'$2'"' >> $T/s
+	echo	export $1 >> $T/s
+	echo fi >> $T/s
+	echo $1_DEF='$'$1  >> $T/s
+	. $T/s
+}
+
+initparam LKMM_DESTDIR "."
+initparam LKMM_HERD_OPTIONS "-conf linux-kernel.cfg"
+initparam LKMM_HW_MAP_FILE ""
+initparam LKMM_JOBS `getconf _NPROCESSORS_ONLN`
+initparam LKMM_PROCS "3"
+initparam LKMM_TIMEOUT "1m"
+
+scriptname=$0
+
+usagehelp () {
+	echo "Usage $scriptname [ arguments ]"
+	echo "      --destdir path (place for .litmus.out, default by .litmus)"
+	echo "      --herdopts -conf linux-kernel.cfg ..."
+	echo "      --hw AArch64"
+	echo "      --jobs N (number of jobs, default one per CPU)"
+	echo "      --procs N (litmus tests with at most this many processes)"
+	echo "      --timeout N (herd7 timeout (e.g., 10s, 1m, 2hr, 1d, '')"
+	echo "Defaults: --destdir '$LKMM_DESTDIR_DEF' --herdopts '$LKMM_HERD_OPTIONS_DEF' --hw '$LKMM_HW_MAP_FILE' --jobs '$LKMM_JOBS_DEF' --procs '$LKMM_PROCS_DEF' --timeout '$LKMM_TIMEOUT_DEF'"
+	exit 1
+}
+
+usage () {
+	usagehelp 1>&2
+}
+
+# checkarg --argname argtype $# arg mustmatch cannotmatch
+checkarg () {
+	if test $3 -le 1
+	then
+		echo $1 needs argument $2 matching \"$5\"
+		usage
+	fi
+	if echo "$4" | grep -q -e "$5"
+	then
+		:
+	else
+		echo $1 $2 \"$4\" must match \"$5\"
+		usage
+	fi
+	if echo "$4" | grep -q -e "$6"
+	then
+		echo $1 $2 \"$4\" must not match \"$6\"
+		usage
+	fi
+}
+
+while test $# -gt 0
+do
+	case "$1" in
+	--destdir)
+		checkarg --destdir "(path to directory)" "$#" "$2" '.\+' '^--'
+		LKMM_DESTDIR="$2"
+		mkdir $LKMM_DESTDIR > /dev/null 2>&1
+		if ! test -e "$LKMM_DESTDIR"
+		then
+			echo "Cannot create directory --destdir '$LKMM_DESTDIR'"
+			usage
+		fi
+		if test -d "$LKMM_DESTDIR" -a -x "$LKMM_DESTDIR"
+		then
+			:
+		else
+			echo "Directory --destdir '$LKMM_DESTDIR' insufficient permissions to create files"
+			usage
+		fi
+		shift
+		;;
+	--herdopts|--herdopt)
+		checkarg --destdir "(herd7 options)" "$#" "$2" '.*' '^--'
+		LKMM_HERD_OPTIONS="$2"
+		shift
+		;;
+	--hw)
+		checkarg --hw "(.map file architecture name)" "$#" "$2" '^[A-Za-z0-9_-]\+' '^--'
+		LKMM_HW_MAP_FILE="$2"
+		shift
+		;;
+	-j[1-9]*)
+		njobs="`echo $1 | sed -e 's/^-j//'`"
+		trailchars="`echo $njobs | sed -e 's/[0-9]\+\(.*\)$/\1/'`"
+		if test -n "$trailchars"
+		then
+			echo $1 trailing characters "'$trailchars'"
+			usagehelp
+		fi
+		LKMM_JOBS="`echo $njobs | sed -e 's/^\([0-9]\+\).*$/\1/'`"
+		;;
+	--jobs|--job|-j)
+		checkarg --jobs "(number)" "$#" "$2" '^[1-9][0-9]*$' '^--'
+		LKMM_JOBS="$2"
+		shift
+		;;
+	--procs|--proc)
+		checkarg --procs "(number)" "$#" "$2" '^[0-9]\+$' '^--'
+		LKMM_PROCS="$2"
+		shift
+		;;
+	--timeout)
+		checkarg --timeout "(timeout spec)" "$#" "$2" '^\([0-9]\+[smhd]\?\|\)$' '^--'
+		LKMM_TIMEOUT="$2"
+		shift
+		;;
+	--)
+		shift
+		break
+		;;
+	*)
+		echo Unknown argument $1
+		usage
+		;;
+	esac
+	shift
+done
+if test -z "$LKMM_TIMEOUT"
+then
+	LKMM_TIMEOUT_CMD=""; export LKMM_TIMEOUT_CMD
+else
+	LKMM_TIMEOUT_CMD="timeout $LKMM_TIMEOUT"; export LKMM_TIMEOUT_CMD
+fi
+rm -rf $T
diff --git a/tools/memory-model/scripts/runlitmus.sh b/tools/memory-model/scripts/runlitmus.sh
new file mode 100755
index 0000000000..94608d4b65
--- /dev/null
+++ b/tools/memory-model/scripts/runlitmus.sh
@@ -0,0 +1,80 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Without the -hw argument, runs a herd7 test and outputs verification
+# results to a file whose name is that of the specified litmus test,
+# but with ".out" appended.
+#
+# If the --hw argument is specified, this script translates the .litmus
+# C-language file to the specified type of assembly and verifies that.
+# But in this case, litmus tests using complex synchronization (such as
+# locking, RCU, and SRCU) are cheerfully ignored.
+#
+# Either way, return the status of the herd7 command.
+#
+# Usage:
+#	runlitmus.sh file.litmus
+#
+# Run this in the directory containing the memory model, specifying the
+# pathname of the litmus test to check.  The caller is expected to have
+# properly set up the LKMM environment variables.
+#
+# Copyright IBM Corporation, 2019
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+litmus=$1
+if test -f "$litmus" -a -r "$litmus"
+then
+	:
+else
+	echo ' !!! ' error: \"$litmus\" is not a readable file
+	exit 255
+fi
+
+if test -z "$LKMM_HW_MAP_FILE" -o ! -e $LKMM_DESTDIR/$litmus.out
+then
+	# LKMM run
+	herdoptions=${LKMM_HERD_OPTIONS--conf linux-kernel.cfg}
+	echo Herd options: $herdoptions > $LKMM_DESTDIR/$litmus.out
+	/usr/bin/time $LKMM_TIMEOUT_CMD herd7 $herdoptions $litmus >> $LKMM_DESTDIR/$litmus.out 2>&1
+	ret=$?
+	if test -z "$LKMM_HW_MAP_FILE"
+	then
+		exit $ret
+	fi
+	echo " --- " Automatically generated LKMM output for '"'--hw $LKMM_HW_MAP_FILE'"' run
+fi
+
+# Hardware run
+
+T=/tmp/checklitmushw.sh.$$
+trap 'rm -rf $T' 0 2
+mkdir $T
+
+# Generate filenames
+mapfile="Linux2${LKMM_HW_MAP_FILE}.map"
+themefile="$T/${LKMM_HW_MAP_FILE}.theme"
+herdoptions="-model $LKMM_HW_CAT_FILE"
+hwlitmus=`echo $litmus | sed -e 's/\.litmus$/.litmus.'${LKMM_HW_MAP_FILE}'/'`
+hwlitmusfile=`echo $hwlitmus | sed -e 's,^.*/,,'`
+
+# Don't run on litmus tests with complex synchronization
+if ! scripts/simpletest.sh $litmus
+then
+	echo ' --- ' error: \"$litmus\" contains locking, RCU, or SRCU
+	exit 254
+fi
+
+# Generate the assembly code and run herd7 on it.
+gen_theme7 -n 10 -map $mapfile -call Linux.call > $themefile
+jingle7 -v -theme $themefile $litmus > $LKMM_DESTDIR/$hwlitmus 2> $T/$hwlitmusfile.jingle7.out
+if grep -q "Generated 0 tests" $T/$hwlitmusfile.jingle7.out
+then
+	echo ' !!! ' jingle7 failed, errors in $hwlitmus.err
+	cp $T/$hwlitmusfile.jingle7.out $LKMM_DESTDIR/$hwlitmus.err
+	exit 253
+fi
+/usr/bin/time $LKMM_TIMEOUT_CMD herd7 -unroll 0 $LKMM_DESTDIR/$hwlitmus > $LKMM_DESTDIR/$hwlitmus.out 2>&1
+
+exit $?
diff --git a/tools/memory-model/scripts/runlitmushist.sh b/tools/memory-model/scripts/runlitmushist.sh
new file mode 100755
index 0000000000..c6c2bdc67a
--- /dev/null
+++ b/tools/memory-model/scripts/runlitmushist.sh
@@ -0,0 +1,94 @@
+#!/bin/bash
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Runs the C-language litmus tests specified on standard input, using up
+# to the specified number of CPUs (defaulting to all of them) and placing
+# the results in the specified directory (defaulting to the same place
+# the litmus test came from).
+#
+# sh runlitmushist.sh
+#
+# Run from the Linux kernel tools/memory-model directory.
+# This script uses environment variables produced by parseargs.sh.
+#
+# Copyright IBM Corporation, 2018
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+. scripts/hwfnseg.sh
+
+T=/tmp/runlitmushist.sh.$$
+trap 'rm -rf $T' 0
+mkdir $T
+
+if test -d litmus
+then
+	:
+else
+	echo Directory \"litmus\" missing, aborting run.
+	exit 1
+fi
+
+# Prefixes for per-CPU scripts
+for ((i=0;i<$LKMM_JOBS;i++))
+do
+	echo T=$T >> $T/$i.sh
+	cat << '___EOF___' >> $T/$i.sh
+	runtest () {
+		if scripts/runlitmus.sh $1
+		then
+			if ! grep -q '^Observation ' $LKMM_DESTDIR/$1$2.out
+			then
+				echo ' !!! Herd failed, no Observation:' $1
+			fi
+		else
+			exitcode=$?
+			if test "$exitcode" -eq 124
+			then
+				exitmsg="timed out"
+			elif test "$exitcode" -eq 253
+			then
+				exitmsg=
+			else
+				exitmsg="failed, exit code $exitcode"
+			fi
+			if test -n "$exitmsg"
+			then
+				echo ' !!! Herd' ${exitmsg}: $1
+			fi
+		fi
+	}
+___EOF___
+done
+
+awk -v q="'" -v b='\\' '
+{
+	print "echo `grep " q "^P[0-9]" b "+(" q " " $0 " | tail -1 | sed -e " q "s/^P" b "([0-9]" b "+" b ")(.*$/" b "1/" q "` " $0
+}' | sh | sort -k1n |
+awk -v dq='"' -v hwfnseg="$hwfnseg" -v ncpu="$LKMM_JOBS" -v t="$T" '
+{
+	print "if test -z " dq hwfnseg dq " || scripts/simpletest.sh " dq $2 dq
+	print "then"
+	print "\techo runtest " dq $2 dq " " hwfnseg " >> " t "/" NR % ncpu ".sh";
+	print "fi"
+}
+
+END {
+	for (i = 0; i < ncpu; i++) {
+		print "sh " t "/" i ".sh > " t "/" i ".sh.out 2>&1 &";
+		close(t "/" i ".sh");
+	}
+	print "wait";
+}' | sh
+cat $T/*.sh.out
+if grep -q '!!!' $T/*.sh.out
+then
+	echo ' ---' Summary: 1>&2
+	grep '!!!' $T/*.sh.out 1>&2
+	nfail="`grep '!!!' $T/*.sh.out | wc -l`"
+	echo 'Number of failed herd7 runs (e.g., timeout): ' $nfail 1>&2
+	exit 1
+else
+	echo All runs completed successfully. 1>&2
+	exit 0
+fi
diff --git a/tools/memory-model/scripts/simpletest.sh b/tools/memory-model/scripts/simpletest.sh
new file mode 100755
index 0000000000..7edc5d3616
--- /dev/null
+++ b/tools/memory-model/scripts/simpletest.sh
@@ -0,0 +1,35 @@
+#!/bin/sh
+# SPDX-License-Identifier: GPL-2.0+
+#
+# Give zero status if this is a simple test and non-zero otherwise.
+# Simple tests do not contain locking, RCU, or SRCU.
+#
+# Usage:
+#	simpletest.sh file.litmus
+#
+# Copyright IBM Corporation, 2019
+#
+# Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+
+litmus=$1
+
+if test -f "$litmus" -a -r "$litmus"
+then
+	:
+else
+	echo ' --- ' error: \"$litmus\" is not a readable file
+	exit 255
+fi
+exclude="^[[:space:]]*\("
+exclude="${exclude}spin_lock(\|spin_unlock(\|spin_trylock(\|spin_is_locked("
+exclude="${exclude}\|rcu_read_lock(\|rcu_read_unlock("
+exclude="${exclude}\|synchronize_rcu(\|synchronize_rcu_expedited("
+exclude="${exclude}\|srcu_read_lock(\|srcu_read_unlock("
+exclude="${exclude}\|synchronize_srcu(\|synchronize_srcu_expedited("
+exclude="${exclude}\)"
+if grep -q $exclude $litmus
+then
+	exit 255
+fi
+exit 0