Adding upstream version 5.10.209.upstream/5.10.209upstream

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

48 files changed, 12277 insertions, 0 deletions

diff --git a/Documentation/core-api/assoc_array.rst b/Documentation/core-api/assoc_array.rst
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--- /dev/null
+++ b/Documentation/core-api/assoc_array.rst
@@ -0,0 +1,554 @@
+========================================
+Generic Associative Array Implementation
+========================================
+
+Overview
+========
+
+This associative array implementation is an object container with the following
+properties:
+
+1. Objects are opaque pointers.  The implementation does not care where they
+   point (if anywhere) or what they point to (if anything).
+
+   .. note::
+
+      Pointers to objects _must_ be zero in the least significant bit.
+
+2. Objects do not need to contain linkage blocks for use by the array.  This
+   permits an object to be located in multiple arrays simultaneously.
+   Rather, the array is made up of metadata blocks that point to objects.
+
+3. Objects require index keys to locate them within the array.
+
+4. Index keys must be unique.  Inserting an object with the same key as one
+   already in the array will replace the old object.
+
+5. Index keys can be of any length and can be of different lengths.
+
+6. Index keys should encode the length early on, before any variation due to
+   length is seen.
+
+7. Index keys can include a hash to scatter objects throughout the array.
+
+8. The array can iterated over.  The objects will not necessarily come out in
+   key order.
+
+9. The array can be iterated over while it is being modified, provided the
+   RCU readlock is being held by the iterator.  Note, however, under these
+   circumstances, some objects may be seen more than once.  If this is a
+   problem, the iterator should lock against modification.  Objects will not
+   be missed, however, unless deleted.
+
+10. Objects in the array can be looked up by means of their index key.
+
+11. Objects can be looked up while the array is being modified, provided the
+    RCU readlock is being held by the thread doing the look up.
+
+The implementation uses a tree of 16-pointer nodes internally that are indexed
+on each level by nibbles from the index key in the same manner as in a radix
+tree.  To improve memory efficiency, shortcuts can be emplaced to skip over
+what would otherwise be a series of single-occupancy nodes.  Further, nodes
+pack leaf object pointers into spare space in the node rather than making an
+extra branch until as such time an object needs to be added to a full node.
+
+
+The Public API
+==============
+
+The public API can be found in ``<linux/assoc_array.h>``.  The associative
+array is rooted on the following structure::
+
+    struct assoc_array {
+            ...
+    };
+
+The code is selected by enabling ``CONFIG_ASSOCIATIVE_ARRAY`` with::
+
+    ./script/config -e ASSOCIATIVE_ARRAY
+
+
+Edit Script
+-----------
+
+The insertion and deletion functions produce an 'edit script' that can later be
+applied to effect the changes without risking ``ENOMEM``. This retains the
+preallocated metadata blocks that will be installed in the internal tree and
+keeps track of the metadata blocks that will be removed from the tree when the
+script is applied.
+
+This is also used to keep track of dead blocks and dead objects after the
+script has been applied so that they can be freed later.  The freeing is done
+after an RCU grace period has passed - thus allowing access functions to
+proceed under the RCU read lock.
+
+The script appears as outside of the API as a pointer of the type::
+
+    struct assoc_array_edit;
+
+There are two functions for dealing with the script:
+
+1. Apply an edit script::
+
+    void assoc_array_apply_edit(struct assoc_array_edit *edit);
+
+This will perform the edit functions, interpolating various write barriers
+to permit accesses under the RCU read lock to continue.  The edit script
+will then be passed to ``call_rcu()`` to free it and any dead stuff it points
+to.
+
+2. Cancel an edit script::
+
+    void assoc_array_cancel_edit(struct assoc_array_edit *edit);
+
+This frees the edit script and all preallocated memory immediately. If
+this was for insertion, the new object is _not_ released by this function,
+but must rather be released by the caller.
+
+These functions are guaranteed not to fail.
+
+
+Operations Table
+----------------
+
+Various functions take a table of operations::
+
+    struct assoc_array_ops {
+            ...
+    };
+
+This points to a number of methods, all of which need to be provided:
+
+1. Get a chunk of index key from caller data::
+
+    unsigned long (*get_key_chunk)(const void *index_key, int level);
+
+This should return a chunk of caller-supplied index key starting at the
+*bit* position given by the level argument.  The level argument will be a
+multiple of ``ASSOC_ARRAY_KEY_CHUNK_SIZE`` and the function should return
+``ASSOC_ARRAY_KEY_CHUNK_SIZE bits``.  No error is possible.
+
+
+2. Get a chunk of an object's index key::
+
+    unsigned long (*get_object_key_chunk)(const void *object, int level);
+
+As the previous function, but gets its data from an object in the array
+rather than from a caller-supplied index key.
+
+
+3. See if this is the object we're looking for::
+
+    bool (*compare_object)(const void *object, const void *index_key);
+
+Compare the object against an index key and return ``true`` if it matches and
+``false`` if it doesn't.
+
+
+4. Diff the index keys of two objects::
+
+    int (*diff_objects)(const void *object, const void *index_key);
+
+Return the bit position at which the index key of the specified object
+differs from the given index key or -1 if they are the same.
+
+
+5. Free an object::
+
+    void (*free_object)(void *object);
+
+Free the specified object.  Note that this may be called an RCU grace period
+after ``assoc_array_apply_edit()`` was called, so ``synchronize_rcu()`` may be
+necessary on module unloading.
+
+
+Manipulation Functions
+----------------------
+
+There are a number of functions for manipulating an associative array:
+
+1. Initialise an associative array::
+
+    void assoc_array_init(struct assoc_array *array);
+
+This initialises the base structure for an associative array.  It can't fail.
+
+
+2. Insert/replace an object in an associative array::
+
+    struct assoc_array_edit *
+    assoc_array_insert(struct assoc_array *array,
+                       const struct assoc_array_ops *ops,
+                       const void *index_key,
+                       void *object);
+
+This inserts the given object into the array.  Note that the least
+significant bit of the pointer must be zero as it's used to type-mark
+pointers internally.
+
+If an object already exists for that key then it will be replaced with the
+new object and the old one will be freed automatically.
+
+The ``index_key`` argument should hold index key information and is
+passed to the methods in the ops table when they are called.
+
+This function makes no alteration to the array itself, but rather returns
+an edit script that must be applied.  ``-ENOMEM`` is returned in the case of
+an out-of-memory error.
+
+The caller should lock exclusively against other modifiers of the array.
+
+
+3. Delete an object from an associative array::
+
+    struct assoc_array_edit *
+    assoc_array_delete(struct assoc_array *array,
+                       const struct assoc_array_ops *ops,
+                       const void *index_key);
+
+This deletes an object that matches the specified data from the array.
+
+The ``index_key`` argument should hold index key information and is
+passed to the methods in the ops table when they are called.
+
+This function makes no alteration to the array itself, but rather returns
+an edit script that must be applied.  ``-ENOMEM`` is returned in the case of
+an out-of-memory error.  ``NULL`` will be returned if the specified object is
+not found within the array.
+
+The caller should lock exclusively against other modifiers of the array.
+
+
+4. Delete all objects from an associative array::
+
+    struct assoc_array_edit *
+    assoc_array_clear(struct assoc_array *array,
+                      const struct assoc_array_ops *ops);
+
+This deletes all the objects from an associative array and leaves it
+completely empty.
+
+This function makes no alteration to the array itself, but rather returns
+an edit script that must be applied.  ``-ENOMEM`` is returned in the case of
+an out-of-memory error.
+
+The caller should lock exclusively against other modifiers of the array.
+
+
+5. Destroy an associative array, deleting all objects::
+
+    void assoc_array_destroy(struct assoc_array *array,
+                             const struct assoc_array_ops *ops);
+
+This destroys the contents of the associative array and leaves it
+completely empty.  It is not permitted for another thread to be traversing
+the array under the RCU read lock at the same time as this function is
+destroying it as no RCU deferral is performed on memory release -
+something that would require memory to be allocated.
+
+The caller should lock exclusively against other modifiers and accessors
+of the array.
+
+
+6. Garbage collect an associative array::
+
+    int assoc_array_gc(struct assoc_array *array,
+                       const struct assoc_array_ops *ops,
+                       bool (*iterator)(void *object, void *iterator_data),
+                       void *iterator_data);
+
+This iterates over the objects in an associative array and passes each one to
+``iterator()``.  If ``iterator()`` returns ``true``, the object is kept.  If it
+returns ``false``, the object will be freed.  If the ``iterator()`` function
+returns ``true``, it must perform any appropriate refcount incrementing on the
+object before returning.
+
+The internal tree will be packed down if possible as part of the iteration
+to reduce the number of nodes in it.
+
+The ``iterator_data`` is passed directly to ``iterator()`` and is otherwise
+ignored by the function.
+
+The function will return ``0`` if successful and ``-ENOMEM`` if there wasn't
+enough memory.
+
+It is possible for other threads to iterate over or search the array under
+the RCU read lock while this function is in progress.  The caller should
+lock exclusively against other modifiers of the array.
+
+
+Access Functions
+----------------
+
+There are two functions for accessing an associative array:
+
+1. Iterate over all the objects in an associative array::
+
+    int assoc_array_iterate(const struct assoc_array *array,
+                            int (*iterator)(const void *object,
+                                            void *iterator_data),
+                            void *iterator_data);
+
+This passes each object in the array to the iterator callback function.
+``iterator_data`` is private data for that function.
+
+This may be used on an array at the same time as the array is being
+modified, provided the RCU read lock is held.  Under such circumstances,
+it is possible for the iteration function to see some objects twice.  If
+this is a problem, then modification should be locked against.  The
+iteration algorithm should not, however, miss any objects.
+
+The function will return ``0`` if no objects were in the array or else it will
+return the result of the last iterator function called.  Iteration stops
+immediately if any call to the iteration function results in a non-zero
+return.
+
+
+2. Find an object in an associative array::
+
+    void *assoc_array_find(const struct assoc_array *array,
+                           const struct assoc_array_ops *ops,
+                           const void *index_key);
+
+This walks through the array's internal tree directly to the object
+specified by the index key..
+
+This may be used on an array at the same time as the array is being
+modified, provided the RCU read lock is held.
+
+The function will return the object if found (and set ``*_type`` to the object
+type) or will return ``NULL`` if the object was not found.
+
+
+Index Key Form
+--------------
+
+The index key can be of any form, but since the algorithms aren't told how long
+the key is, it is strongly recommended that the index key includes its length
+very early on before any variation due to the length would have an effect on
+comparisons.
+
+This will cause leaves with different length keys to scatter away from each
+other - and those with the same length keys to cluster together.
+
+It is also recommended that the index key begin with a hash of the rest of the
+key to maximise scattering throughout keyspace.
+
+The better the scattering, the wider and lower the internal tree will be.
+
+Poor scattering isn't too much of a problem as there are shortcuts and nodes
+can contain mixtures of leaves and metadata pointers.
+
+The index key is read in chunks of machine word.  Each chunk is subdivided into
+one nibble (4 bits) per level, so on a 32-bit CPU this is good for 8 levels and
+on a 64-bit CPU, 16 levels.  Unless the scattering is really poor, it is
+unlikely that more than one word of any particular index key will have to be
+used.
+
+
+Internal Workings
+=================
+
+The associative array data structure has an internal tree.  This tree is
+constructed of two types of metadata blocks: nodes and shortcuts.
+
+A node is an array of slots.  Each slot can contain one of four things:
+
+* A NULL pointer, indicating that the slot is empty.
+* A pointer to an object (a leaf).
+* A pointer to a node at the next level.
+* A pointer to a shortcut.
+
+
+Basic Internal Tree Layout
+--------------------------
+
+Ignoring shortcuts for the moment, the nodes form a multilevel tree.  The index
+key space is strictly subdivided by the nodes in the tree and nodes occur on
+fixed levels.  For example::
+
+ Level: 0               1               2               3
+        =============== =============== =============== ===============
+                                                        NODE D
+                        NODE B          NODE C  +------>+---+
+                +------>+---+   +------>+---+   |       | 0 |
+        NODE A  |       | 0 |   |       | 0 |   |       +---+
+        +---+   |       +---+   |       +---+   |       :   :
+        | 0 |   |       :   :   |       :   :   |       +---+
+        +---+   |       +---+   |       +---+   |       | f |
+        | 1 |---+       | 3 |---+       | 7 |---+       +---+
+        +---+           +---+           +---+
+        :   :           :   :           | 8 |---+
+        +---+           +---+           +---+   |       NODE E
+        | e |---+       | f |           :   :   +------>+---+
+        +---+   |       +---+           +---+           | 0 |
+        | f |   |                       | f |           +---+
+        +---+   |                       +---+           :   :
+                |       NODE F                          +---+
+                +------>+---+                           | f |
+                        | 0 |           NODE G          +---+
+                        +---+   +------>+---+
+                        :   :   |       | 0 |
+                        +---+   |       +---+
+                        | 6 |---+       :   :
+                        +---+           +---+
+                        :   :           | f |
+                        +---+           +---+
+                        | f |
+                        +---+
+
+In the above example, there are 7 nodes (A-G), each with 16 slots (0-f).
+Assuming no other meta data nodes in the tree, the key space is divided
+thusly::
+
+    KEY PREFIX      NODE
+    ==========      ====
+    137*            D
+    138*            E
+    13[0-69-f]*     C
+    1[0-24-f]*      B
+    e6*             G
+    e[0-57-f]*      F
+    [02-df]*        A
+
+So, for instance, keys with the following example index keys will be found in
+the appropriate nodes::
+
+    INDEX KEY       PREFIX  NODE
+    =============== ======= ====
+    13694892892489  13      C
+    13795289025897  137     D
+    13889dde88793   138     E
+    138bbb89003093  138     E
+    1394879524789   12      C
+    1458952489      1       B
+    9431809de993ba  -       A
+    b4542910809cd   -       A
+    e5284310def98   e       F
+    e68428974237    e6      G
+    e7fffcbd443     e       F
+    f3842239082     -       A
+
+To save memory, if a node can hold all the leaves in its portion of keyspace,
+then the node will have all those leaves in it and will not have any metadata
+pointers - even if some of those leaves would like to be in the same slot.
+
+A node can contain a heterogeneous mix of leaves and metadata pointers.
+Metadata pointers must be in the slots that match their subdivisions of key
+space.  The leaves can be in any slot not occupied by a metadata pointer.  It
+is guaranteed that none of the leaves in a node will match a slot occupied by a
+metadata pointer.  If the metadata pointer is there, any leaf whose key matches
+the metadata key prefix must be in the subtree that the metadata pointer points
+to.
+
+In the above example list of index keys, node A will contain::
+
+    SLOT    CONTENT         INDEX KEY (PREFIX)
+    ====    =============== ==================
+    1       PTR TO NODE B   1*
+    any     LEAF            9431809de993ba
+    any     LEAF            b4542910809cd
+    e       PTR TO NODE F   e*
+    any     LEAF            f3842239082
+
+and node B::
+
+    3	PTR TO NODE C	13*
+    any	LEAF		1458952489
+
+
+Shortcuts
+---------
+
+Shortcuts are metadata records that jump over a piece of keyspace.  A shortcut
+is a replacement for a series of single-occupancy nodes ascending through the
+levels.  Shortcuts exist to save memory and to speed up traversal.
+
+It is possible for the root of the tree to be a shortcut - say, for example,
+the tree contains at least 17 nodes all with key prefix ``1111``.  The
+insertion algorithm will insert a shortcut to skip over the ``1111`` keyspace
+in a single bound and get to the fourth level where these actually become
+different.
+
+
+Splitting And Collapsing Nodes
+------------------------------
+
+Each node has a maximum capacity of 16 leaves and metadata pointers.  If the
+insertion algorithm finds that it is trying to insert a 17th object into a
+node, that node will be split such that at least two leaves that have a common
+key segment at that level end up in a separate node rooted on that slot for
+that common key segment.
+
+If the leaves in a full node and the leaf that is being inserted are
+sufficiently similar, then a shortcut will be inserted into the tree.
+
+When the number of objects in the subtree rooted at a node falls to 16 or
+fewer, then the subtree will be collapsed down to a single node - and this will
+ripple towards the root if possible.
+
+
+Non-Recursive Iteration
+-----------------------
+
+Each node and shortcut contains a back pointer to its parent and the number of
+slot in that parent that points to it.  None-recursive iteration uses these to
+proceed rootwards through the tree, going to the parent node, slot N + 1 to
+make sure progress is made without the need for a stack.
+
+The backpointers, however, make simultaneous alteration and iteration tricky.
+
+
+Simultaneous Alteration And Iteration
+-------------------------------------
+
+There are a number of cases to consider:
+
+1. Simple insert/replace.  This involves simply replacing a NULL or old
+   matching leaf pointer with the pointer to the new leaf after a barrier.
+   The metadata blocks don't change otherwise.  An old leaf won't be freed
+   until after the RCU grace period.
+
+2. Simple delete.  This involves just clearing an old matching leaf.  The
+   metadata blocks don't change otherwise.  The old leaf won't be freed until
+   after the RCU grace period.
+
+3. Insertion replacing part of a subtree that we haven't yet entered.  This
+   may involve replacement of part of that subtree - but that won't affect
+   the iteration as we won't have reached the pointer to it yet and the
+   ancestry blocks are not replaced (the layout of those does not change).
+
+4. Insertion replacing nodes that we're actively processing.  This isn't a
+   problem as we've passed the anchoring pointer and won't switch onto the
+   new layout until we follow the back pointers - at which point we've
+   already examined the leaves in the replaced node (we iterate over all the
+   leaves in a node before following any of its metadata pointers).
+
+   We might, however, re-see some leaves that have been split out into a new
+   branch that's in a slot further along than we were at.
+
+5. Insertion replacing nodes that we're processing a dependent branch of.
+   This won't affect us until we follow the back pointers.  Similar to (4).
+
+6. Deletion collapsing a branch under us.  This doesn't affect us because the
+   back pointers will get us back to the parent of the new node before we
+   could see the new node.  The entire collapsed subtree is thrown away
+   unchanged - and will still be rooted on the same slot, so we shouldn't
+   process it a second time as we'll go back to slot + 1.
+
+.. note::
+
+   Under some circumstances, we need to simultaneously change the parent
+   pointer and the parent slot pointer on a node (say, for example, we
+   inserted another node before it and moved it up a level).  We cannot do
+   this without locking against a read - so we have to replace that node too.
+
+   However, when we're changing a shortcut into a node this isn't a problem
+   as shortcuts only have one slot and so the parent slot number isn't used
+   when traversing backwards over one.  This means that it's okay to change
+   the slot number first - provided suitable barriers are used to make sure
+   the parent slot number is read after the back pointer.
+
+Obsolete blocks and leaves are freed up after an RCU grace period has passed,
+so as long as anyone doing walking or iteration holds the RCU read lock, the
+old superstructure should not go away on them.
diff --git a/Documentation/core-api/atomic_ops.rst b/Documentation/core-api/atomic_ops.rst
new file mode 100644
index 000000000..724583453
--- /dev/null
+++ b/Documentation/core-api/atomic_ops.rst
@@ -0,0 +1,664 @@
+=======================================================
+Semantics and Behavior of Atomic and Bitmask Operations
+=======================================================
+
+:Author: David S. Miller
+
+This document is intended to serve as a guide to Linux port
+maintainers on how to implement atomic counter, bitops, and spinlock
+interfaces properly.
+
+Atomic Type And Operations
+==========================
+
+The atomic_t type should be defined as a signed integer and
+the atomic_long_t type as a signed long integer.  Also, they should
+be made opaque such that any kind of cast to a normal C integer type
+will fail.  Something like the following should suffice::
+
+	typedef struct { int counter; } atomic_t;
+	typedef struct { long counter; } atomic_long_t;
+
+Historically, counter has been declared volatile.  This is now discouraged.
+See :ref:`Documentation/process/volatile-considered-harmful.rst
+<volatile_considered_harmful>` for the complete rationale.
+
+local_t is very similar to atomic_t. If the counter is per CPU and only
+updated by one CPU, local_t is probably more appropriate. Please see
+:ref:`Documentation/core-api/local_ops.rst <local_ops>` for the semantics of
+local_t.
+
+The first operations to implement for atomic_t's are the initializers and
+plain writes. ::
+
+	#define ATOMIC_INIT(i)		{ (i) }
+	#define atomic_set(v, i)	((v)->counter = (i))
+
+The first macro is used in definitions, such as::
+
+	static atomic_t my_counter = ATOMIC_INIT(1);
+
+The initializer is atomic in that the return values of the atomic operations
+are guaranteed to be correct reflecting the initialized value if the
+initializer is used before runtime.  If the initializer is used at runtime, a
+proper implicit or explicit read memory barrier is needed before reading the
+value with atomic_read from another thread.
+
+As with all of the ``atomic_`` interfaces, replace the leading ``atomic_``
+with ``atomic_long_`` to operate on atomic_long_t.
+
+The second interface can be used at runtime, as in::
+
+	struct foo { atomic_t counter; };
+	...
+
+	struct foo *k;
+
+	k = kmalloc(sizeof(*k), GFP_KERNEL);
+	if (!k)
+		return -ENOMEM;
+	atomic_set(&k->counter, 0);
+
+The setting is atomic in that the return values of the atomic operations by
+all threads are guaranteed to be correct reflecting either the value that has
+been set with this operation or set with another operation.  A proper implicit
+or explicit memory barrier is needed before the value set with the operation
+is guaranteed to be readable with atomic_read from another thread.
+
+Next, we have::
+
+	#define atomic_read(v)	((v)->counter)
+
+which simply reads the counter value currently visible to the calling thread.
+The read is atomic in that the return value is guaranteed to be one of the
+values initialized or modified with the interface operations if a proper
+implicit or explicit memory barrier is used after possible runtime
+initialization by any other thread and the value is modified only with the
+interface operations.  atomic_read does not guarantee that the runtime
+initialization by any other thread is visible yet, so the user of the
+interface must take care of that with a proper implicit or explicit memory
+barrier.
+
+.. warning::
+
+	``atomic_read()`` and ``atomic_set()`` DO NOT IMPLY BARRIERS!
+
+	Some architectures may choose to use the volatile keyword, barriers, or
+	inline assembly to guarantee some degree of immediacy for atomic_read()
+	and atomic_set().  This is not uniformly guaranteed, and may change in
+	the future, so all users of atomic_t should treat atomic_read() and
+	atomic_set() as simple C statements that may be reordered or optimized
+	away entirely by the compiler or processor, and explicitly invoke the
+	appropriate compiler and/or memory barrier for each use case.  Failure
+	to do so will result in code that may suddenly break when used with
+	different architectures or compiler optimizations, or even changes in
+	unrelated code which changes how the compiler optimizes the section
+	accessing atomic_t variables.
+
+Properly aligned pointers, longs, ints, and chars (and unsigned
+equivalents) may be atomically loaded from and stored to in the same
+sense as described for atomic_read() and atomic_set().  The READ_ONCE()
+and WRITE_ONCE() macros should be used to prevent the compiler from using
+optimizations that might otherwise optimize accesses out of existence on
+the one hand, or that might create unsolicited accesses on the other.
+
+For example consider the following code::
+
+	while (a > 0)
+		do_something();
+
+If the compiler can prove that do_something() does not store to the
+variable a, then the compiler is within its rights transforming this to
+the following::
+
+	if (a > 0)
+		for (;;)
+			do_something();
+
+If you don't want the compiler to do this (and you probably don't), then
+you should use something like the following::
+
+	while (READ_ONCE(a) > 0)
+		do_something();
+
+Alternatively, you could place a barrier() call in the loop.
+
+For another example, consider the following code::
+
+	tmp_a = a;
+	do_something_with(tmp_a);
+	do_something_else_with(tmp_a);
+
+If the compiler can prove that do_something_with() does not store to the
+variable a, then the compiler is within its rights to manufacture an
+additional load as follows::
+
+	tmp_a = a;
+	do_something_with(tmp_a);
+	tmp_a = a;
+	do_something_else_with(tmp_a);
+
+This could fatally confuse your code if it expected the same value
+to be passed to do_something_with() and do_something_else_with().
+
+The compiler would be likely to manufacture this additional load if
+do_something_with() was an inline function that made very heavy use
+of registers: reloading from variable a could save a flush to the
+stack and later reload.  To prevent the compiler from attacking your
+code in this manner, write the following::
+
+	tmp_a = READ_ONCE(a);
+	do_something_with(tmp_a);
+	do_something_else_with(tmp_a);
+
+For a final example, consider the following code, assuming that the
+variable a is set at boot time before the second CPU is brought online
+and never changed later, so that memory barriers are not needed::
+
+	if (a)
+		b = 9;
+	else
+		b = 42;
+
+The compiler is within its rights to manufacture an additional store
+by transforming the above code into the following::
+
+	b = 42;
+	if (a)
+		b = 9;
+
+This could come as a fatal surprise to other code running concurrently
+that expected b to never have the value 42 if a was zero.  To prevent
+the compiler from doing this, write something like::
+
+	if (a)
+		WRITE_ONCE(b, 9);
+	else
+		WRITE_ONCE(b, 42);
+
+Don't even -think- about doing this without proper use of memory barriers,
+locks, or atomic operations if variable a can change at runtime!
+
+.. warning::
+
+	``READ_ONCE()`` OR ``WRITE_ONCE()`` DO NOT IMPLY A BARRIER!
+
+Now, we move onto the atomic operation interfaces typically implemented with
+the help of assembly code. ::
+
+	void atomic_add(int i, atomic_t *v);
+	void atomic_sub(int i, atomic_t *v);
+	void atomic_inc(atomic_t *v);
+	void atomic_dec(atomic_t *v);
+
+These four routines add and subtract integral values to/from the given
+atomic_t value.  The first two routines pass explicit integers by
+which to make the adjustment, whereas the latter two use an implicit
+adjustment value of "1".
+
+One very important aspect of these two routines is that they DO NOT
+require any explicit memory barriers.  They need only perform the
+atomic_t counter update in an SMP safe manner.
+
+Next, we have::
+
+	int atomic_inc_return(atomic_t *v);
+	int atomic_dec_return(atomic_t *v);
+
+These routines add 1 and subtract 1, respectively, from the given
+atomic_t and return the new counter value after the operation is
+performed.
+
+Unlike the above routines, it is required that these primitives
+include explicit memory barriers that are performed before and after
+the operation.  It must be done such that all memory operations before
+and after the atomic operation calls are strongly ordered with respect
+to the atomic operation itself.
+
+For example, it should behave as if a smp_mb() call existed both
+before and after the atomic operation.
+
+If the atomic instructions used in an implementation provide explicit
+memory barrier semantics which satisfy the above requirements, that is
+fine as well.
+
+Let's move on::
+
+	int atomic_add_return(int i, atomic_t *v);
+	int atomic_sub_return(int i, atomic_t *v);
+
+These behave just like atomic_{inc,dec}_return() except that an
+explicit counter adjustment is given instead of the implicit "1".
+This means that like atomic_{inc,dec}_return(), the memory barrier
+semantics are required.
+
+Next::
+
+	int atomic_inc_and_test(atomic_t *v);
+	int atomic_dec_and_test(atomic_t *v);
+
+These two routines increment and decrement by 1, respectively, the
+given atomic counter.  They return a boolean indicating whether the
+resulting counter value was zero or not.
+
+Again, these primitives provide explicit memory barrier semantics around
+the atomic operation::
+
+	int atomic_sub_and_test(int i, atomic_t *v);
+
+This is identical to atomic_dec_and_test() except that an explicit
+decrement is given instead of the implicit "1".  This primitive must
+provide explicit memory barrier semantics around the operation::
+
+	int atomic_add_negative(int i, atomic_t *v);
+
+The given increment is added to the given atomic counter value.  A boolean
+is return which indicates whether the resulting counter value is negative.
+This primitive must provide explicit memory barrier semantics around
+the operation.
+
+Then::
+
+	int atomic_xchg(atomic_t *v, int new);
+
+This performs an atomic exchange operation on the atomic variable v, setting
+the given new value.  It returns the old value that the atomic variable v had
+just before the operation.
+
+atomic_xchg must provide explicit memory barriers around the operation. ::
+
+	int atomic_cmpxchg(atomic_t *v, int old, int new);
+
+This performs an atomic compare exchange operation on the atomic value v,
+with the given old and new values. Like all atomic_xxx operations,
+atomic_cmpxchg will only satisfy its atomicity semantics as long as all
+other accesses of \*v are performed through atomic_xxx operations.
+
+atomic_cmpxchg must provide explicit memory barriers around the operation,
+although if the comparison fails then no memory ordering guarantees are
+required.
+
+The semantics for atomic_cmpxchg are the same as those defined for 'cas'
+below.
+
+Finally::
+
+	int atomic_add_unless(atomic_t *v, int a, int u);
+
+If the atomic value v is not equal to u, this function adds a to v, and
+returns non zero. If v is equal to u then it returns zero. This is done as
+an atomic operation.
+
+atomic_add_unless must provide explicit memory barriers around the
+operation unless it fails (returns 0).
+
+atomic_inc_not_zero, equivalent to atomic_add_unless(v, 1, 0)
+
+
+If a caller requires memory barrier semantics around an atomic_t
+operation which does not return a value, a set of interfaces are
+defined which accomplish this::
+
+	void smp_mb__before_atomic(void);
+	void smp_mb__after_atomic(void);
+
+Preceding a non-value-returning read-modify-write atomic operation with
+smp_mb__before_atomic() and following it with smp_mb__after_atomic()
+provides the same full ordering that is provided by value-returning
+read-modify-write atomic operations.
+
+For example, smp_mb__before_atomic() can be used like so::
+
+	obj->dead = 1;
+	smp_mb__before_atomic();
+	atomic_dec(&obj->ref_count);
+
+It makes sure that all memory operations preceding the atomic_dec()
+call are strongly ordered with respect to the atomic counter
+operation.  In the above example, it guarantees that the assignment of
+"1" to obj->dead will be globally visible to other cpus before the
+atomic counter decrement.
+
+Without the explicit smp_mb__before_atomic() call, the
+implementation could legally allow the atomic counter update visible
+to other cpus before the "obj->dead = 1;" assignment.
+
+A missing memory barrier in the cases where they are required by the
+atomic_t implementation above can have disastrous results.  Here is
+an example, which follows a pattern occurring frequently in the Linux
+kernel.  It is the use of atomic counters to implement reference
+counting, and it works such that once the counter falls to zero it can
+be guaranteed that no other entity can be accessing the object::
+
+	static void obj_list_add(struct obj *obj, struct list_head *head)
+	{
+		obj->active = 1;
+		list_add(&obj->list, head);
+	}
+
+	static void obj_list_del(struct obj *obj)
+	{
+		list_del(&obj->list);
+		obj->active = 0;
+	}
+
+	static void obj_destroy(struct obj *obj)
+	{
+		BUG_ON(obj->active);
+		kfree(obj);
+	}
+
+	struct obj *obj_list_peek(struct list_head *head)
+	{
+		if (!list_empty(head)) {
+			struct obj *obj;
+
+			obj = list_entry(head->next, struct obj, list);
+			atomic_inc(&obj->refcnt);
+			return obj;
+		}
+		return NULL;
+	}
+
+	void obj_poke(void)
+	{
+		struct obj *obj;
+
+		spin_lock(&global_list_lock);
+		obj = obj_list_peek(&global_list);
+		spin_unlock(&global_list_lock);
+
+		if (obj) {
+			obj->ops->poke(obj);
+			if (atomic_dec_and_test(&obj->refcnt))
+				obj_destroy(obj);
+		}
+	}
+
+	void obj_timeout(struct obj *obj)
+	{
+		spin_lock(&global_list_lock);
+		obj_list_del(obj);
+		spin_unlock(&global_list_lock);
+
+		if (atomic_dec_and_test(&obj->refcnt))
+			obj_destroy(obj);
+	}
+
+.. note::
+
+	This is a simplification of the ARP queue management in the generic
+	neighbour discover code of the networking.  Olaf Kirch found a bug wrt.
+	memory barriers in kfree_skb() that exposed the atomic_t memory barrier
+	requirements quite clearly.
+
+Given the above scheme, it must be the case that the obj->active
+update done by the obj list deletion be visible to other processors
+before the atomic counter decrement is performed.
+
+Otherwise, the counter could fall to zero, yet obj->active would still
+be set, thus triggering the assertion in obj_destroy().  The error
+sequence looks like this::
+
+	cpu 0				cpu 1
+	obj_poke()			obj_timeout()
+	obj = obj_list_peek();
+	... gains ref to obj, refcnt=2
+					obj_list_del(obj);
+					obj->active = 0 ...
+					... visibility delayed ...
+					atomic_dec_and_test()
+					... refcnt drops to 1 ...
+	atomic_dec_and_test()
+	... refcount drops to 0 ...
+	obj_destroy()
+	BUG() triggers since obj->active
+	still seen as one
+					obj->active update visibility occurs
+
+With the memory barrier semantics required of the atomic_t operations
+which return values, the above sequence of memory visibility can never
+happen.  Specifically, in the above case the atomic_dec_and_test()
+counter decrement would not become globally visible until the
+obj->active update does.
+
+As a historical note, 32-bit Sparc used to only allow usage of
+24-bits of its atomic_t type.  This was because it used 8 bits
+as a spinlock for SMP safety.  Sparc32 lacked a "compare and swap"
+type instruction.  However, 32-bit Sparc has since been moved over
+to a "hash table of spinlocks" scheme, that allows the full 32-bit
+counter to be realized.  Essentially, an array of spinlocks are
+indexed into based upon the address of the atomic_t being operated
+on, and that lock protects the atomic operation.  Parisc uses the
+same scheme.
+
+Another note is that the atomic_t operations returning values are
+extremely slow on an old 386.
+
+
+Atomic Bitmask
+==============
+
+We will now cover the atomic bitmask operations.  You will find that
+their SMP and memory barrier semantics are similar in shape and scope
+to the atomic_t ops above.
+
+Native atomic bit operations are defined to operate on objects aligned
+to the size of an "unsigned long" C data type, and are least of that
+size.  The endianness of the bits within each "unsigned long" are the
+native endianness of the cpu. ::
+
+	void set_bit(unsigned long nr, volatile unsigned long *addr);
+	void clear_bit(unsigned long nr, volatile unsigned long *addr);
+	void change_bit(unsigned long nr, volatile unsigned long *addr);
+
+These routines set, clear, and change, respectively, the bit number
+indicated by "nr" on the bit mask pointed to by "ADDR".
+
+They must execute atomically, yet there are no implicit memory barrier
+semantics required of these interfaces. ::
+
+	int test_and_set_bit(unsigned long nr, volatile unsigned long *addr);
+	int test_and_clear_bit(unsigned long nr, volatile unsigned long *addr);
+	int test_and_change_bit(unsigned long nr, volatile unsigned long *addr);
+
+Like the above, except that these routines return a boolean which
+indicates whether the changed bit was set _BEFORE_ the atomic bit
+operation.
+
+
+.. warning::
+        It is incredibly important that the value be a boolean, ie. "0" or "1".
+        Do not try to be fancy and save a few instructions by declaring the
+        above to return "long" and just returning something like "old_val &
+        mask" because that will not work.
+
+For one thing, this return value gets truncated to int in many code
+paths using these interfaces, so on 64-bit if the bit is set in the
+upper 32-bits then testers will never see that.
+
+One great example of where this problem crops up are the thread_info
+flag operations.  Routines such as test_and_set_ti_thread_flag() chop
+the return value into an int.  There are other places where things
+like this occur as well.
+
+These routines, like the atomic_t counter operations returning values,
+must provide explicit memory barrier semantics around their execution.
+All memory operations before the atomic bit operation call must be
+made visible globally before the atomic bit operation is made visible.
+Likewise, the atomic bit operation must be visible globally before any
+subsequent memory operation is made visible.  For example::
+
+	obj->dead = 1;
+	if (test_and_set_bit(0, &obj->flags))
+		/* ... */;
+	obj->killed = 1;
+
+The implementation of test_and_set_bit() must guarantee that
+"obj->dead = 1;" is visible to cpus before the atomic memory operation
+done by test_and_set_bit() becomes visible.  Likewise, the atomic
+memory operation done by test_and_set_bit() must become visible before
+"obj->killed = 1;" is visible.
+
+Finally there is the basic operation::
+
+	int test_bit(unsigned long nr, __const__ volatile unsigned long *addr);
+
+Which returns a boolean indicating if bit "nr" is set in the bitmask
+pointed to by "addr".
+
+If explicit memory barriers are required around {set,clear}_bit() (which do
+not return a value, and thus does not need to provide memory barrier
+semantics), two interfaces are provided::
+
+	void smp_mb__before_atomic(void);
+	void smp_mb__after_atomic(void);
+
+They are used as follows, and are akin to their atomic_t operation
+brothers::
+
+	/* All memory operations before this call will
+	 * be globally visible before the clear_bit().
+	 */
+	smp_mb__before_atomic();
+	clear_bit( ... );
+
+	/* The clear_bit() will be visible before all
+	 * subsequent memory operations.
+	 */
+	 smp_mb__after_atomic();
+
+There are two special bitops with lock barrier semantics (acquire/release,
+same as spinlocks). These operate in the same way as their non-_lock/unlock
+postfixed variants, except that they are to provide acquire/release semantics,
+respectively. This means they can be used for bit_spin_trylock and
+bit_spin_unlock type operations without specifying any more barriers. ::
+
+	int test_and_set_bit_lock(unsigned long nr, unsigned long *addr);
+	void clear_bit_unlock(unsigned long nr, unsigned long *addr);
+	void __clear_bit_unlock(unsigned long nr, unsigned long *addr);
+
+The __clear_bit_unlock version is non-atomic, however it still implements
+unlock barrier semantics. This can be useful if the lock itself is protecting
+the other bits in the word.
+
+Finally, there are non-atomic versions of the bitmask operations
+provided.  They are used in contexts where some other higher-level SMP
+locking scheme is being used to protect the bitmask, and thus less
+expensive non-atomic operations may be used in the implementation.
+They have names similar to the above bitmask operation interfaces,
+except that two underscores are prefixed to the interface name. ::
+
+	void __set_bit(unsigned long nr, volatile unsigned long *addr);
+	void __clear_bit(unsigned long nr, volatile unsigned long *addr);
+	void __change_bit(unsigned long nr, volatile unsigned long *addr);
+	int __test_and_set_bit(unsigned long nr, volatile unsigned long *addr);
+	int __test_and_clear_bit(unsigned long nr, volatile unsigned long *addr);
+	int __test_and_change_bit(unsigned long nr, volatile unsigned long *addr);
+
+These non-atomic variants also do not require any special memory
+barrier semantics.
+
+The routines xchg() and cmpxchg() must provide the same exact
+memory-barrier semantics as the atomic and bit operations returning
+values.
+
+.. note::
+
+	If someone wants to use xchg(), cmpxchg() and their variants,
+	linux/atomic.h should be included rather than asm/cmpxchg.h, unless the
+	code is in arch/* and can take care of itself.
+
+Spinlocks and rwlocks have memory barrier expectations as well.
+The rule to follow is simple:
+
+1) When acquiring a lock, the implementation must make it globally
+   visible before any subsequent memory operation.
+
+2) When releasing a lock, the implementation must make it such that
+   all previous memory operations are globally visible before the
+   lock release.
+
+Which finally brings us to _atomic_dec_and_lock().  There is an
+architecture-neutral version implemented in lib/dec_and_lock.c,
+but most platforms will wish to optimize this in assembler. ::
+
+	int _atomic_dec_and_lock(atomic_t *atomic, spinlock_t *lock);
+
+Atomically decrement the given counter, and if will drop to zero
+atomically acquire the given spinlock and perform the decrement
+of the counter to zero.  If it does not drop to zero, do nothing
+with the spinlock.
+
+It is actually pretty simple to get the memory barrier correct.
+Simply satisfy the spinlock grab requirements, which is make
+sure the spinlock operation is globally visible before any
+subsequent memory operation.
+
+We can demonstrate this operation more clearly if we define
+an abstract atomic operation::
+
+	long cas(long *mem, long old, long new);
+
+"cas" stands for "compare and swap".  It atomically:
+
+1) Compares "old" with the value currently at "mem".
+2) If they are equal, "new" is written to "mem".
+3) Regardless, the current value at "mem" is returned.
+
+As an example usage, here is what an atomic counter update
+might look like::
+
+	void example_atomic_inc(long *counter)
+	{
+		long old, new, ret;
+
+		while (1) {
+			old = *counter;
+			new = old + 1;
+
+			ret = cas(counter, old, new);
+			if (ret == old)
+				break;
+		}
+	}
+
+Let's use cas() in order to build a pseudo-C atomic_dec_and_lock()::
+
+	int _atomic_dec_and_lock(atomic_t *atomic, spinlock_t *lock)
+	{
+		long old, new, ret;
+		int went_to_zero;
+
+		went_to_zero = 0;
+		while (1) {
+			old = atomic_read(atomic);
+			new = old - 1;
+			if (new == 0) {
+				went_to_zero = 1;
+				spin_lock(lock);
+			}
+			ret = cas(atomic, old, new);
+			if (ret == old)
+				break;
+			if (went_to_zero) {
+				spin_unlock(lock);
+				went_to_zero = 0;
+			}
+		}
+
+		return went_to_zero;
+	}
+
+Now, as far as memory barriers go, as long as spin_lock()
+strictly orders all subsequent memory operations (including
+the cas()) with respect to itself, things will be fine.
+
+Said another way, _atomic_dec_and_lock() must guarantee that
+a counter dropping to zero is never made visible before the
+spinlock being acquired.
+
+.. note::
+
+	Note that this also means that for the case where the counter is not
+	dropping to zero, there are no memory ordering requirements.
diff --git a/Documentation/core-api/boot-time-mm.rst b/Documentation/core-api/boot-time-mm.rst
new file mode 100644
index 000000000..e5ec9f1a5
--- /dev/null
+++ b/Documentation/core-api/boot-time-mm.rst
@@ -0,0 +1,41 @@
+===========================
+Boot time memory management
+===========================
+
+Early system initialization cannot use "normal" memory management
+simply because it is not set up yet. But there is still need to
+allocate memory for various data structures, for instance for the
+physical page allocator.
+
+A specialized allocator called ``memblock`` performs the
+boot time memory management. The architecture specific initialization
+must set it up in :c:func:`setup_arch` and tear it down in
+:c:func:`mem_init` functions.
+
+Once the early memory management is available it offers a variety of
+functions and macros for memory allocations. The allocation request
+may be directed to the first (and probably the only) node or to a
+particular node in a NUMA system. There are API variants that panic
+when an allocation fails and those that don't.
+
+Memblock also offers a variety of APIs that control its own behaviour.
+
+Memblock Overview
+=================
+
+.. kernel-doc:: mm/memblock.c
+   :doc: memblock overview
+
+
+Functions and structures
+========================
+
+Here is the description of memblock data structures, functions and
+macros. Some of them are actually internal, but since they are
+documented it would be silly to omit them. Besides, reading the
+descriptions for the internal functions can help to understand what
+really happens under the hood.
+
+.. kernel-doc:: include/linux/memblock.h
+.. kernel-doc:: mm/memblock.c
+   :functions:
diff --git a/Documentation/core-api/bus-virt-phys-mapping.rst b/Documentation/core-api/bus-virt-phys-mapping.rst
new file mode 100644
index 000000000..c7bc99cd2
--- /dev/null
+++ b/Documentation/core-api/bus-virt-phys-mapping.rst
@@ -0,0 +1,220 @@
+==========================================================
+How to access I/O mapped memory from within device drivers
+==========================================================
+
+:Author: Linus
+
+.. warning::
+
+	The virt_to_bus() and bus_to_virt() functions have been
+	superseded by the functionality provided by the PCI DMA interface
+	(see :doc:`/core-api/dma-api-howto`).  They continue
+	to be documented below for historical purposes, but new code
+	must not use them. --davidm 00/12/12
+
+::
+
+  [ This is a mail message in response to a query on IO mapping, thus the
+    strange format for a "document" ]
+
+The AHA-1542 is a bus-master device, and your patch makes the driver give the
+controller the physical address of the buffers, which is correct on x86
+(because all bus master devices see the physical memory mappings directly). 
+
+However, on many setups, there are actually **three** different ways of looking
+at memory addresses, and in this case we actually want the third, the
+so-called "bus address". 
+
+Essentially, the three ways of addressing memory are (this is "real memory",
+that is, normal RAM--see later about other details): 
+
+ - CPU untranslated.  This is the "physical" address.  Physical address 
+   0 is what the CPU sees when it drives zeroes on the memory bus.
+
+ - CPU translated address. This is the "virtual" address, and is 
+   completely internal to the CPU itself with the CPU doing the appropriate
+   translations into "CPU untranslated". 
+
+ - bus address. This is the address of memory as seen by OTHER devices, 
+   not the CPU. Now, in theory there could be many different bus 
+   addresses, with each device seeing memory in some device-specific way, but
+   happily most hardware designers aren't actually actively trying to make
+   things any more complex than necessary, so you can assume that all 
+   external hardware sees the memory the same way. 
+
+Now, on normal PCs the bus address is exactly the same as the physical
+address, and things are very simple indeed. However, they are that simple
+because the memory and the devices share the same address space, and that is
+not generally necessarily true on other PCI/ISA setups. 
+
+Now, just as an example, on the PReP (PowerPC Reference Platform), the 
+CPU sees a memory map something like this (this is from memory)::
+
+	0-2 GB		"real memory"
+	2 GB-3 GB	"system IO" (inb/out and similar accesses on x86)
+	3 GB-4 GB 	"IO memory" (shared memory over the IO bus)
+
+Now, that looks simple enough. However, when you look at the same thing from
+the viewpoint of the devices, you have the reverse, and the physical memory
+address 0 actually shows up as address 2 GB for any IO master.
+
+So when the CPU wants any bus master to write to physical memory 0, it 
+has to give the master address 0x80000000 as the memory address.
+
+So, for example, depending on how the kernel is actually mapped on the 
+PPC, you can end up with a setup like this::
+
+ physical address:	0
+ virtual address:	0xC0000000
+ bus address:		0x80000000
+
+where all the addresses actually point to the same thing.  It's just seen 
+through different translations..
+
+Similarly, on the Alpha, the normal translation is::
+
+ physical address:	0
+ virtual address:	0xfffffc0000000000
+ bus address:		0x40000000
+
+(but there are also Alphas where the physical address and the bus address
+are the same). 
+
+Anyway, the way to look up all these translations, you do::
+
+	#include <asm/io.h>
+
+	phys_addr = virt_to_phys(virt_addr);
+	virt_addr = phys_to_virt(phys_addr);
+	 bus_addr = virt_to_bus(virt_addr);
+	virt_addr = bus_to_virt(bus_addr);
+
+Now, when do you need these?
+
+You want the **virtual** address when you are actually going to access that
+pointer from the kernel. So you can have something like this::
+
+	/*
+	 * this is the hardware "mailbox" we use to communicate with
+	 * the controller. The controller sees this directly.
+	 */
+	struct mailbox {
+		__u32 status;
+		__u32 bufstart;
+		__u32 buflen;
+		..
+	} mbox;
+
+		unsigned char * retbuffer;
+
+		/* get the address from the controller */
+		retbuffer = bus_to_virt(mbox.bufstart);
+		switch (retbuffer[0]) {
+			case STATUS_OK:
+				...
+
+on the other hand, you want the bus address when you have a buffer that 
+you want to give to the controller::
+
+	/* ask the controller to read the sense status into "sense_buffer" */
+	mbox.bufstart = virt_to_bus(&sense_buffer);
+	mbox.buflen = sizeof(sense_buffer);
+	mbox.status = 0;
+	notify_controller(&mbox);
+
+And you generally **never** want to use the physical address, because you can't
+use that from the CPU (the CPU only uses translated virtual addresses), and
+you can't use it from the bus master. 
+
+So why do we care about the physical address at all? We do need the physical
+address in some cases, it's just not very often in normal code.  The physical
+address is needed if you use memory mappings, for example, because the
+"remap_pfn_range()" mm function wants the physical address of the memory to
+be remapped as measured in units of pages, a.k.a. the pfn (the memory
+management layer doesn't know about devices outside the CPU, so it
+shouldn't need to know about "bus addresses" etc).
+
+.. note::
+
+	The above is only one part of the whole equation. The above
+	only talks about "real memory", that is, CPU memory (RAM).
+
+There is a completely different type of memory too, and that's the "shared
+memory" on the PCI or ISA bus. That's generally not RAM (although in the case
+of a video graphics card it can be normal DRAM that is just used for a frame
+buffer), but can be things like a packet buffer in a network card etc. 
+
+This memory is called "PCI memory" or "shared memory" or "IO memory" or
+whatever, and there is only one way to access it: the readb/writeb and
+related functions. You should never take the address of such memory, because
+there is really nothing you can do with such an address: it's not
+conceptually in the same memory space as "real memory" at all, so you cannot
+just dereference a pointer. (Sadly, on x86 it **is** in the same memory space,
+so on x86 it actually works to just deference a pointer, but it's not
+portable). 
+
+For such memory, you can do things like:
+
+ - reading::
+
+	/*
+	 * read first 32 bits from ISA memory at 0xC0000, aka
+	 * C000:0000 in DOS terms
+	 */
+	unsigned int signature = isa_readl(0xC0000);
+
+ - remapping and writing::
+
+	/*
+	 * remap framebuffer PCI memory area at 0xFC000000,
+	 * size 1MB, so that we can access it: We can directly
+	 * access only the 640k-1MB area, so anything else
+	 * has to be remapped.
+	 */
+	void __iomem *baseptr = ioremap(0xFC000000, 1024*1024);
+
+	/* write a 'A' to the offset 10 of the area */
+	writeb('A',baseptr+10);
+
+	/* unmap when we unload the driver */
+	iounmap(baseptr);
+
+ - copying and clearing::
+
+	/* get the 6-byte Ethernet address at ISA address E000:0040 */
+	memcpy_fromio(kernel_buffer, 0xE0040, 6);
+	/* write a packet to the driver */
+	memcpy_toio(0xE1000, skb->data, skb->len);
+	/* clear the frame buffer */
+	memset_io(0xA0000, 0, 0x10000);
+
+OK, that just about covers the basics of accessing IO portably.  Questions?
+Comments? You may think that all the above is overly complex, but one day you
+might find yourself with a 500 MHz Alpha in front of you, and then you'll be
+happy that your driver works ;)
+
+Note that kernel versions 2.0.x (and earlier) mistakenly called the
+ioremap() function "vremap()".  ioremap() is the proper name, but I
+didn't think straight when I wrote it originally.  People who have to
+support both can do something like::
+ 
+	/* support old naming silliness */
+	#if LINUX_VERSION_CODE < 0x020100
+	#define ioremap vremap
+	#define iounmap vfree                                                     
+	#endif
+ 
+at the top of their source files, and then they can use the right names
+even on 2.0.x systems. 
+
+And the above sounds worse than it really is.  Most real drivers really
+don't do all that complex things (or rather: the complexity is not so
+much in the actual IO accesses as in error handling and timeouts etc). 
+It's generally not hard to fix drivers, and in many cases the code
+actually looks better afterwards::
+
+	unsigned long signature = *(unsigned int *) 0xC0000;
+		vs
+	unsigned long signature = readl(0xC0000);
+
+I think the second version actually is more readable, no?
diff --git a/Documentation/core-api/cachetlb.rst b/Documentation/core-api/cachetlb.rst
new file mode 100644
index 000000000..a1582cc79
--- /dev/null
+++ b/Documentation/core-api/cachetlb.rst
@@ -0,0 +1,405 @@
+==================================
+Cache and TLB Flushing Under Linux
+==================================
+
+:Author: David S. Miller <davem@redhat.com>
+
+This document describes the cache/tlb flushing interfaces called
+by the Linux VM subsystem.  It enumerates over each interface,
+describes its intended purpose, and what side effect is expected
+after the interface is invoked.
+
+The side effects described below are stated for a uniprocessor
+implementation, and what is to happen on that single processor.  The
+SMP cases are a simple extension, in that you just extend the
+definition such that the side effect for a particular interface occurs
+on all processors in the system.  Don't let this scare you into
+thinking SMP cache/tlb flushing must be so inefficient, this is in
+fact an area where many optimizations are possible.  For example,
+if it can be proven that a user address space has never executed
+on a cpu (see mm_cpumask()), one need not perform a flush
+for this address space on that cpu.
+
+First, the TLB flushing interfaces, since they are the simplest.  The
+"TLB" is abstracted under Linux as something the cpu uses to cache
+virtual-->physical address translations obtained from the software
+page tables.  Meaning that if the software page tables change, it is
+possible for stale translations to exist in this "TLB" cache.
+Therefore when software page table changes occur, the kernel will
+invoke one of the following flush methods _after_ the page table
+changes occur:
+
+1) ``void flush_tlb_all(void)``
+
+	The most severe flush of all.  After this interface runs,
+	any previous page table modification whatsoever will be
+	visible to the cpu.
+
+	This is usually invoked when the kernel page tables are
+	changed, since such translations are "global" in nature.
+
+2) ``void flush_tlb_mm(struct mm_struct *mm)``
+
+	This interface flushes an entire user address space from
+	the TLB.  After running, this interface must make sure that
+	any previous page table modifications for the address space
+	'mm' will be visible to the cpu.  That is, after running,
+	there will be no entries in the TLB for 'mm'.
+
+	This interface is used to handle whole address space
+	page table operations such as what happens during
+	fork, and exec.
+
+3) ``void flush_tlb_range(struct vm_area_struct *vma,
+   unsigned long start, unsigned long end)``
+
+	Here we are flushing a specific range of (user) virtual
+	address translations from the TLB.  After running, this
+	interface must make sure that any previous page table
+	modifications for the address space 'vma->vm_mm' in the range
+	'start' to 'end-1' will be visible to the cpu.  That is, after
+	running, there will be no entries in the TLB for 'mm' for
+	virtual addresses in the range 'start' to 'end-1'.
+
+	The "vma" is the backing store being used for the region.
+	Primarily, this is used for munmap() type operations.
+
+	The interface is provided in hopes that the port can find
+	a suitably efficient method for removing multiple page
+	sized translations from the TLB, instead of having the kernel
+	call flush_tlb_page (see below) for each entry which may be
+	modified.
+
+4) ``void flush_tlb_page(struct vm_area_struct *vma, unsigned long addr)``
+
+	This time we need to remove the PAGE_SIZE sized translation
+	from the TLB.  The 'vma' is the backing structure used by
+	Linux to keep track of mmap'd regions for a process, the
+	address space is available via vma->vm_mm.  Also, one may
+	test (vma->vm_flags & VM_EXEC) to see if this region is
+	executable (and thus could be in the 'instruction TLB' in
+	split-tlb type setups).
+
+	After running, this interface must make sure that any previous
+	page table modification for address space 'vma->vm_mm' for
+	user virtual address 'addr' will be visible to the cpu.  That
+	is, after running, there will be no entries in the TLB for
+	'vma->vm_mm' for virtual address 'addr'.
+
+	This is used primarily during fault processing.
+
+5) ``void update_mmu_cache(struct vm_area_struct *vma,
+   unsigned long address, pte_t *ptep)``
+
+	At the end of every page fault, this routine is invoked to
+	tell the architecture specific code that a translation
+	now exists at virtual address "address" for address space
+	"vma->vm_mm", in the software page tables.
+
+	A port may use this information in any way it so chooses.
+	For example, it could use this event to pre-load TLB
+	translations for software managed TLB configurations.
+	The sparc64 port currently does this.
+
+Next, we have the cache flushing interfaces.  In general, when Linux
+is changing an existing virtual-->physical mapping to a new value,
+the sequence will be in one of the following forms::
+
+	1) flush_cache_mm(mm);
+	   change_all_page_tables_of(mm);
+	   flush_tlb_mm(mm);
+
+	2) flush_cache_range(vma, start, end);
+	   change_range_of_page_tables(mm, start, end);
+	   flush_tlb_range(vma, start, end);
+
+	3) flush_cache_page(vma, addr, pfn);
+	   set_pte(pte_pointer, new_pte_val);
+	   flush_tlb_page(vma, addr);
+
+The cache level flush will always be first, because this allows
+us to properly handle systems whose caches are strict and require
+a virtual-->physical translation to exist for a virtual address
+when that virtual address is flushed from the cache.  The HyperSparc
+cpu is one such cpu with this attribute.
+
+The cache flushing routines below need only deal with cache flushing
+to the extent that it is necessary for a particular cpu.  Mostly,
+these routines must be implemented for cpus which have virtually
+indexed caches which must be flushed when virtual-->physical
+translations are changed or removed.  So, for example, the physically
+indexed physically tagged caches of IA32 processors have no need to
+implement these interfaces since the caches are fully synchronized
+and have no dependency on translation information.
+
+Here are the routines, one by one:
+
+1) ``void flush_cache_mm(struct mm_struct *mm)``
+
+	This interface flushes an entire user address space from
+	the caches.  That is, after running, there will be no cache
+	lines associated with 'mm'.
+
+	This interface is used to handle whole address space
+	page table operations such as what happens during exit and exec.
+
+2) ``void flush_cache_dup_mm(struct mm_struct *mm)``
+
+	This interface flushes an entire user address space from
+	the caches.  That is, after running, there will be no cache
+	lines associated with 'mm'.
+
+	This interface is used to handle whole address space
+	page table operations such as what happens during fork.
+
+	This option is separate from flush_cache_mm to allow some
+	optimizations for VIPT caches.
+
+3) ``void flush_cache_range(struct vm_area_struct *vma,
+   unsigned long start, unsigned long end)``
+
+	Here we are flushing a specific range of (user) virtual
+	addresses from the cache.  After running, there will be no
+	entries in the cache for 'vma->vm_mm' for virtual addresses in
+	the range 'start' to 'end-1'.
+
+	The "vma" is the backing store being used for the region.
+	Primarily, this is used for munmap() type operations.
+
+	The interface is provided in hopes that the port can find
+	a suitably efficient method for removing multiple page
+	sized regions from the cache, instead of having the kernel
+	call flush_cache_page (see below) for each entry which may be
+	modified.
+
+4) ``void flush_cache_page(struct vm_area_struct *vma, unsigned long addr, unsigned long pfn)``
+
+	This time we need to remove a PAGE_SIZE sized range
+	from the cache.  The 'vma' is the backing structure used by
+	Linux to keep track of mmap'd regions for a process, the
+	address space is available via vma->vm_mm.  Also, one may
+	test (vma->vm_flags & VM_EXEC) to see if this region is
+	executable (and thus could be in the 'instruction cache' in
+	"Harvard" type cache layouts).
+
+	The 'pfn' indicates the physical page frame (shift this value
+	left by PAGE_SHIFT to get the physical address) that 'addr'
+	translates to.  It is this mapping which should be removed from
+	the cache.
+
+	After running, there will be no entries in the cache for
+	'vma->vm_mm' for virtual address 'addr' which translates
+	to 'pfn'.
+
+	This is used primarily during fault processing.
+
+5) ``void flush_cache_kmaps(void)``
+
+	This routine need only be implemented if the platform utilizes
+	highmem.  It will be called right before all of the kmaps
+	are invalidated.
+
+	After running, there will be no entries in the cache for
+	the kernel virtual address range PKMAP_ADDR(0) to
+	PKMAP_ADDR(LAST_PKMAP).
+
+	This routing should be implemented in asm/highmem.h
+
+6) ``void flush_cache_vmap(unsigned long start, unsigned long end)``
+   ``void flush_cache_vunmap(unsigned long start, unsigned long end)``
+
+	Here in these two interfaces we are flushing a specific range
+	of (kernel) virtual addresses from the cache.  After running,
+	there will be no entries in the cache for the kernel address
+	space for virtual addresses in the range 'start' to 'end-1'.
+
+	The first of these two routines is invoked after map_kernel_range()
+	has installed the page table entries.  The second is invoked
+	before unmap_kernel_range() deletes the page table entries.
+
+There exists another whole class of cpu cache issues which currently
+require a whole different set of interfaces to handle properly.
+The biggest problem is that of virtual aliasing in the data cache
+of a processor.
+
+Is your port susceptible to virtual aliasing in its D-cache?
+Well, if your D-cache is virtually indexed, is larger in size than
+PAGE_SIZE, and does not prevent multiple cache lines for the same
+physical address from existing at once, you have this problem.
+
+If your D-cache has this problem, first define asm/shmparam.h SHMLBA
+properly, it should essentially be the size of your virtually
+addressed D-cache (or if the size is variable, the largest possible
+size).  This setting will force the SYSv IPC layer to only allow user
+processes to mmap shared memory at address which are a multiple of
+this value.
+
+.. note::
+
+  This does not fix shared mmaps, check out the sparc64 port for
+  one way to solve this (in particular SPARC_FLAG_MMAPSHARED).
+
+Next, you have to solve the D-cache aliasing issue for all
+other cases.  Please keep in mind that fact that, for a given page
+mapped into some user address space, there is always at least one more
+mapping, that of the kernel in its linear mapping starting at
+PAGE_OFFSET.  So immediately, once the first user maps a given
+physical page into its address space, by implication the D-cache
+aliasing problem has the potential to exist since the kernel already
+maps this page at its virtual address.
+
+  ``void copy_user_page(void *to, void *from, unsigned long addr, struct page *page)``
+  ``void clear_user_page(void *to, unsigned long addr, struct page *page)``
+
+	These two routines store data in user anonymous or COW
+	pages.  It allows a port to efficiently avoid D-cache alias
+	issues between userspace and the kernel.
+
+	For example, a port may temporarily map 'from' and 'to' to
+	kernel virtual addresses during the copy.  The virtual address
+	for these two pages is chosen in such a way that the kernel
+	load/store instructions happen to virtual addresses which are
+	of the same "color" as the user mapping of the page.  Sparc64
+	for example, uses this technique.
+
+	The 'addr' parameter tells the virtual address where the
+	user will ultimately have this page mapped, and the 'page'
+	parameter gives a pointer to the struct page of the target.
+
+	If D-cache aliasing is not an issue, these two routines may
+	simply call memcpy/memset directly and do nothing more.
+
+  ``void flush_dcache_page(struct page *page)``
+
+	Any time the kernel writes to a page cache page, _OR_
+	the kernel is about to read from a page cache page and
+	user space shared/writable mappings of this page potentially
+	exist, this routine is called.
+
+	.. note::
+
+	      This routine need only be called for page cache pages
+	      which can potentially ever be mapped into the address
+	      space of a user process.  So for example, VFS layer code
+	      handling vfs symlinks in the page cache need not call
+	      this interface at all.
+
+	The phrase "kernel writes to a page cache page" means,
+	specifically, that the kernel executes store instructions
+	that dirty data in that page at the page->virtual mapping
+	of that page.  It is important to flush here to handle
+	D-cache aliasing, to make sure these kernel stores are
+	visible to user space mappings of that page.
+
+	The corollary case is just as important, if there are users
+	which have shared+writable mappings of this file, we must make
+	sure that kernel reads of these pages will see the most recent
+	stores done by the user.
+
+	If D-cache aliasing is not an issue, this routine may
+	simply be defined as a nop on that architecture.
+
+        There is a bit set aside in page->flags (PG_arch_1) as
+	"architecture private".  The kernel guarantees that,
+	for pagecache pages, it will clear this bit when such
+	a page first enters the pagecache.
+
+	This allows these interfaces to be implemented much more
+	efficiently.  It allows one to "defer" (perhaps indefinitely)
+	the actual flush if there are currently no user processes
+	mapping this page.  See sparc64's flush_dcache_page and
+	update_mmu_cache implementations for an example of how to go
+	about doing this.
+
+	The idea is, first at flush_dcache_page() time, if
+	page->mapping->i_mmap is an empty tree, just mark the architecture
+	private page flag bit.  Later, in update_mmu_cache(), a check is
+	made of this flag bit, and if set the flush is done and the flag
+	bit is cleared.
+
+	.. important::
+
+			It is often important, if you defer the flush,
+			that the actual flush occurs on the same CPU
+			as did the cpu stores into the page to make it
+			dirty.  Again, see sparc64 for examples of how
+			to deal with this.
+
+  ``void copy_to_user_page(struct vm_area_struct *vma, struct page *page,
+  unsigned long user_vaddr, void *dst, void *src, int len)``
+  ``void copy_from_user_page(struct vm_area_struct *vma, struct page *page,
+  unsigned long user_vaddr, void *dst, void *src, int len)``
+
+	When the kernel needs to copy arbitrary data in and out
+	of arbitrary user pages (f.e. for ptrace()) it will use
+	these two routines.
+
+	Any necessary cache flushing or other coherency operations
+	that need to occur should happen here.  If the processor's
+	instruction cache does not snoop cpu stores, it is very
+	likely that you will need to flush the instruction cache
+	for copy_to_user_page().
+
+  ``void flush_anon_page(struct vm_area_struct *vma, struct page *page,
+  unsigned long vmaddr)``
+
+  	When the kernel needs to access the contents of an anonymous
+	page, it calls this function (currently only
+	get_user_pages()).  Note: flush_dcache_page() deliberately
+	doesn't work for an anonymous page.  The default
+	implementation is a nop (and should remain so for all coherent
+	architectures).  For incoherent architectures, it should flush
+	the cache of the page at vmaddr.
+
+  ``void flush_kernel_dcache_page(struct page *page)``
+
+	When the kernel needs to modify a user page is has obtained
+	with kmap, it calls this function after all modifications are
+	complete (but before kunmapping it) to bring the underlying
+	page up to date.  It is assumed here that the user has no
+	incoherent cached copies (i.e. the original page was obtained
+	from a mechanism like get_user_pages()).  The default
+	implementation is a nop and should remain so on all coherent
+	architectures.  On incoherent architectures, this should flush
+	the kernel cache for page (using page_address(page)).
+
+
+  ``void flush_icache_range(unsigned long start, unsigned long end)``
+
+  	When the kernel stores into addresses that it will execute
+	out of (eg when loading modules), this function is called.
+
+	If the icache does not snoop stores then this routine will need
+	to flush it.
+
+  ``void flush_icache_page(struct vm_area_struct *vma, struct page *page)``
+
+	All the functionality of flush_icache_page can be implemented in
+	flush_dcache_page and update_mmu_cache. In the future, the hope
+	is to remove this interface completely.
+
+The final category of APIs is for I/O to deliberately aliased address
+ranges inside the kernel.  Such aliases are set up by use of the
+vmap/vmalloc API.  Since kernel I/O goes via physical pages, the I/O
+subsystem assumes that the user mapping and kernel offset mapping are
+the only aliases.  This isn't true for vmap aliases, so anything in
+the kernel trying to do I/O to vmap areas must manually manage
+coherency.  It must do this by flushing the vmap range before doing
+I/O and invalidating it after the I/O returns.
+
+  ``void flush_kernel_vmap_range(void *vaddr, int size)``
+
+       flushes the kernel cache for a given virtual address range in
+       the vmap area.  This is to make sure that any data the kernel
+       modified in the vmap range is made visible to the physical
+       page.  The design is to make this area safe to perform I/O on.
+       Note that this API does *not* also flush the offset map alias
+       of the area.
+
+  ``void invalidate_kernel_vmap_range(void *vaddr, int size) invalidates``
+
+       the cache for a given virtual address range in the vmap area
+       which prevents the processor from making the cache stale by
+       speculatively reading data while the I/O was occurring to the
+       physical pages.  This is only necessary for data reads into the
+       vmap area.
diff --git a/Documentation/core-api/circular-buffers.rst b/Documentation/core-api/circular-buffers.rst
new file mode 100644
index 000000000..50966f66e
--- /dev/null
+++ b/Documentation/core-api/circular-buffers.rst
@@ -0,0 +1,237 @@
+================
+Circular Buffers
+================
+
+:Author: David Howells <dhowells@redhat.com>
+:Author: Paul E. McKenney <paulmck@linux.ibm.com>
+
+
+Linux provides a number of features that can be used to implement circular
+buffering.  There are two sets of such features:
+
+ (1) Convenience functions for determining information about power-of-2 sized
+     buffers.
+
+ (2) Memory barriers for when the producer and the consumer of objects in the
+     buffer don't want to share a lock.
+
+To use these facilities, as discussed below, there needs to be just one
+producer and just one consumer.  It is possible to handle multiple producers by
+serialising them, and to handle multiple consumers by serialising them.
+
+
+.. Contents:
+
+ (*) What is a circular buffer?
+
+ (*) Measuring power-of-2 buffers.
+
+ (*) Using memory barriers with circular buffers.
+     - The producer.
+     - The consumer.
+
+
+
+What is a circular buffer?
+==========================
+
+First of all, what is a circular buffer?  A circular buffer is a buffer of
+fixed, finite size into which there are two indices:
+
+ (1) A 'head' index - the point at which the producer inserts items into the
+     buffer.
+
+ (2) A 'tail' index - the point at which the consumer finds the next item in
+     the buffer.
+
+Typically when the tail pointer is equal to the head pointer, the buffer is
+empty; and the buffer is full when the head pointer is one less than the tail
+pointer.
+
+The head index is incremented when items are added, and the tail index when
+items are removed.  The tail index should never jump the head index, and both
+indices should be wrapped to 0 when they reach the end of the buffer, thus
+allowing an infinite amount of data to flow through the buffer.
+
+Typically, items will all be of the same unit size, but this isn't strictly
+required to use the techniques below.  The indices can be increased by more
+than 1 if multiple items or variable-sized items are to be included in the
+buffer, provided that neither index overtakes the other.  The implementer must
+be careful, however, as a region more than one unit in size may wrap the end of
+the buffer and be broken into two segments.
+
+Measuring power-of-2 buffers
+============================
+
+Calculation of the occupancy or the remaining capacity of an arbitrarily sized
+circular buffer would normally be a slow operation, requiring the use of a
+modulus (divide) instruction.  However, if the buffer is of a power-of-2 size,
+then a much quicker bitwise-AND instruction can be used instead.
+
+Linux provides a set of macros for handling power-of-2 circular buffers.  These
+can be made use of by::
+
+	#include <linux/circ_buf.h>
+
+The macros are:
+
+ (#) Measure the remaining capacity of a buffer::
+
+	CIRC_SPACE(head_index, tail_index, buffer_size);
+
+     This returns the amount of space left in the buffer[1] into which items
+     can be inserted.
+
+
+ (#) Measure the maximum consecutive immediate space in a buffer::
+
+	CIRC_SPACE_TO_END(head_index, tail_index, buffer_size);
+
+     This returns the amount of consecutive space left in the buffer[1] into
+     which items can be immediately inserted without having to wrap back to the
+     beginning of the buffer.
+
+
+ (#) Measure the occupancy of a buffer::
+
+	CIRC_CNT(head_index, tail_index, buffer_size);
+
+     This returns the number of items currently occupying a buffer[2].
+
+
+ (#) Measure the non-wrapping occupancy of a buffer::
+
+	CIRC_CNT_TO_END(head_index, tail_index, buffer_size);
+
+     This returns the number of consecutive items[2] that can be extracted from
+     the buffer without having to wrap back to the beginning of the buffer.
+
+
+Each of these macros will nominally return a value between 0 and buffer_size-1,
+however:
+
+ (1) CIRC_SPACE*() are intended to be used in the producer.  To the producer
+     they will return a lower bound as the producer controls the head index,
+     but the consumer may still be depleting the buffer on another CPU and
+     moving the tail index.
+
+     To the consumer it will show an upper bound as the producer may be busy
+     depleting the space.
+
+ (2) CIRC_CNT*() are intended to be used in the consumer.  To the consumer they
+     will return a lower bound as the consumer controls the tail index, but the
+     producer may still be filling the buffer on another CPU and moving the
+     head index.
+
+     To the producer it will show an upper bound as the consumer may be busy
+     emptying the buffer.
+
+ (3) To a third party, the order in which the writes to the indices by the
+     producer and consumer become visible cannot be guaranteed as they are
+     independent and may be made on different CPUs - so the result in such a
+     situation will merely be a guess, and may even be negative.
+
+Using memory barriers with circular buffers
+===========================================
+
+By using memory barriers in conjunction with circular buffers, you can avoid
+the need to:
+
+ (1) use a single lock to govern access to both ends of the buffer, thus
+     allowing the buffer to be filled and emptied at the same time; and
+
+ (2) use atomic counter operations.
+
+There are two sides to this: the producer that fills the buffer, and the
+consumer that empties it.  Only one thing should be filling a buffer at any one
+time, and only one thing should be emptying a buffer at any one time, but the
+two sides can operate simultaneously.
+
+
+The producer
+------------
+
+The producer will look something like this::
+
+	spin_lock(&producer_lock);
+
+	unsigned long head = buffer->head;
+	/* The spin_unlock() and next spin_lock() provide needed ordering. */
+	unsigned long tail = READ_ONCE(buffer->tail);
+
+	if (CIRC_SPACE(head, tail, buffer->size) >= 1) {
+		/* insert one item into the buffer */
+		struct item *item = buffer[head];
+
+		produce_item(item);
+
+		smp_store_release(buffer->head,
+				  (head + 1) & (buffer->size - 1));
+
+		/* wake_up() will make sure that the head is committed before
+		 * waking anyone up */
+		wake_up(consumer);
+	}
+
+	spin_unlock(&producer_lock);
+
+This will instruct the CPU that the contents of the new item must be written
+before the head index makes it available to the consumer and then instructs the
+CPU that the revised head index must be written before the consumer is woken.
+
+Note that wake_up() does not guarantee any sort of barrier unless something
+is actually awakened.  We therefore cannot rely on it for ordering.  However,
+there is always one element of the array left empty.  Therefore, the
+producer must produce two elements before it could possibly corrupt the
+element currently being read by the consumer.  Therefore, the unlock-lock
+pair between consecutive invocations of the consumer provides the necessary
+ordering between the read of the index indicating that the consumer has
+vacated a given element and the write by the producer to that same element.
+
+
+The Consumer
+------------
+
+The consumer will look something like this::
+
+	spin_lock(&consumer_lock);
+
+	/* Read index before reading contents at that index. */
+	unsigned long head = smp_load_acquire(buffer->head);
+	unsigned long tail = buffer->tail;
+
+	if (CIRC_CNT(head, tail, buffer->size) >= 1) {
+
+		/* extract one item from the buffer */
+		struct item *item = buffer[tail];
+
+		consume_item(item);
+
+		/* Finish reading descriptor before incrementing tail. */
+		smp_store_release(buffer->tail,
+				  (tail + 1) & (buffer->size - 1));
+	}
+
+	spin_unlock(&consumer_lock);
+
+This will instruct the CPU to make sure the index is up to date before reading
+the new item, and then it shall make sure the CPU has finished reading the item
+before it writes the new tail pointer, which will erase the item.
+
+Note the use of READ_ONCE() and smp_load_acquire() to read the
+opposition index.  This prevents the compiler from discarding and
+reloading its cached value.  This isn't strictly needed if you can
+be sure that the opposition index will _only_ be used the once.
+The smp_load_acquire() additionally forces the CPU to order against
+subsequent memory references.  Similarly, smp_store_release() is used
+in both algorithms to write the thread's index.  This documents the
+fact that we are writing to something that can be read concurrently,
+prevents the compiler from tearing the store, and enforces ordering
+against previous accesses.
+
+
+Further reading
+===============
+
+See also Documentation/memory-barriers.txt for a description of Linux's memory
+barrier facilities.
diff --git a/Documentation/core-api/cpu_hotplug.rst b/Documentation/core-api/cpu_hotplug.rst
new file mode 100644
index 000000000..a2c96bec5
--- /dev/null
+++ b/Documentation/core-api/cpu_hotplug.rst
@@ -0,0 +1,365 @@
+=========================
+CPU hotplug in the Kernel
+=========================
+
+:Date: December, 2016
+:Author: Sebastian Andrzej Siewior <bigeasy@linutronix.de>,
+          Rusty Russell <rusty@rustcorp.com.au>,
+          Srivatsa Vaddagiri <vatsa@in.ibm.com>,
+          Ashok Raj <ashok.raj@intel.com>,
+          Joel Schopp <jschopp@austin.ibm.com>
+
+Introduction
+============
+
+Modern advances in system architectures have introduced advanced error
+reporting and correction capabilities in processors. There are couple OEMS that
+support NUMA hardware which are hot pluggable as well, where physical node
+insertion and removal require support for CPU hotplug.
+
+Such advances require CPUs available to a kernel to be removed either for
+provisioning reasons, or for RAS purposes to keep an offending CPU off
+system execution path. Hence the need for CPU hotplug support in the
+Linux kernel.
+
+A more novel use of CPU-hotplug support is its use today in suspend resume
+support for SMP. Dual-core and HT support makes even a laptop run SMP kernels
+which didn't support these methods.
+
+
+Command Line Switches
+=====================
+``maxcpus=n``
+  Restrict boot time CPUs to *n*. Say if you have four CPUs, using
+  ``maxcpus=2`` will only boot two. You can choose to bring the
+  other CPUs later online.
+
+``nr_cpus=n``
+  Restrict the total amount of CPUs the kernel will support. If the number
+  supplied here is lower than the number of physically available CPUs, then
+  those CPUs can not be brought online later.
+
+``additional_cpus=n``
+  Use this to limit hotpluggable CPUs. This option sets
+  ``cpu_possible_mask = cpu_present_mask + additional_cpus``
+
+  This option is limited to the IA64 architecture.
+
+``possible_cpus=n``
+  This option sets ``possible_cpus`` bits in ``cpu_possible_mask``.
+
+  This option is limited to the X86 and S390 architecture.
+
+``cpu0_hotplug``
+  Allow to shutdown CPU0.
+
+  This option is limited to the X86 architecture.
+
+CPU maps
+========
+
+``cpu_possible_mask``
+  Bitmap of possible CPUs that can ever be available in the
+  system. This is used to allocate some boot time memory for per_cpu variables
+  that aren't designed to grow/shrink as CPUs are made available or removed.
+  Once set during boot time discovery phase, the map is static, i.e no bits
+  are added or removed anytime. Trimming it accurately for your system needs
+  upfront can save some boot time memory.
+
+``cpu_online_mask``
+  Bitmap of all CPUs currently online. Its set in ``__cpu_up()``
+  after a CPU is available for kernel scheduling and ready to receive
+  interrupts from devices. Its cleared when a CPU is brought down using
+  ``__cpu_disable()``, before which all OS services including interrupts are
+  migrated to another target CPU.
+
+``cpu_present_mask``
+  Bitmap of CPUs currently present in the system. Not all
+  of them may be online. When physical hotplug is processed by the relevant
+  subsystem (e.g ACPI) can change and new bit either be added or removed
+  from the map depending on the event is hot-add/hot-remove. There are currently
+  no locking rules as of now. Typical usage is to init topology during boot,
+  at which time hotplug is disabled.
+
+You really don't need to manipulate any of the system CPU maps. They should
+be read-only for most use. When setting up per-cpu resources almost always use
+``cpu_possible_mask`` or ``for_each_possible_cpu()`` to iterate. To macro
+``for_each_cpu()`` can be used to iterate over a custom CPU mask.
+
+Never use anything other than ``cpumask_t`` to represent bitmap of CPUs.
+
+
+Using CPU hotplug
+=================
+The kernel option *CONFIG_HOTPLUG_CPU* needs to be enabled. It is currently
+available on multiple architectures including ARM, MIPS, PowerPC and X86. The
+configuration is done via the sysfs interface: ::
+
+ $ ls -lh /sys/devices/system/cpu
+ total 0
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu0
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu1
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu2
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu3
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu4
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu5
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu6
+ drwxr-xr-x  9 root root    0 Dec 21 16:33 cpu7
+ drwxr-xr-x  2 root root    0 Dec 21 16:33 hotplug
+ -r--r--r--  1 root root 4.0K Dec 21 16:33 offline
+ -r--r--r--  1 root root 4.0K Dec 21 16:33 online
+ -r--r--r--  1 root root 4.0K Dec 21 16:33 possible
+ -r--r--r--  1 root root 4.0K Dec 21 16:33 present
+
+The files *offline*, *online*, *possible*, *present* represent the CPU masks.
+Each CPU folder contains an *online* file which controls the logical on (1) and
+off (0) state. To logically shutdown CPU4: ::
+
+ $ echo 0 > /sys/devices/system/cpu/cpu4/online
+  smpboot: CPU 4 is now offline
+
+Once the CPU is shutdown, it will be removed from */proc/interrupts*,
+*/proc/cpuinfo* and should also not be shown visible by the *top* command. To
+bring CPU4 back online: ::
+
+ $ echo 1 > /sys/devices/system/cpu/cpu4/online
+ smpboot: Booting Node 0 Processor 4 APIC 0x1
+
+The CPU is usable again. This should work on all CPUs. CPU0 is often special
+and excluded from CPU hotplug. On X86 the kernel option
+*CONFIG_BOOTPARAM_HOTPLUG_CPU0* has to be enabled in order to be able to
+shutdown CPU0. Alternatively the kernel command option *cpu0_hotplug* can be
+used. Some known dependencies of CPU0:
+
+* Resume from hibernate/suspend. Hibernate/suspend will fail if CPU0 is offline.
+* PIC interrupts. CPU0 can't be removed if a PIC interrupt is detected.
+
+Please let Fenghua Yu <fenghua.yu@intel.com> know if you find any dependencies
+on CPU0.
+
+The CPU hotplug coordination
+============================
+
+The offline case
+----------------
+Once a CPU has been logically shutdown the teardown callbacks of registered
+hotplug states will be invoked, starting with ``CPUHP_ONLINE`` and terminating
+at state ``CPUHP_OFFLINE``. This includes:
+
+* If tasks are frozen due to a suspend operation then *cpuhp_tasks_frozen*
+  will be set to true.
+* All processes are migrated away from this outgoing CPU to new CPUs.
+  The new CPU is chosen from each process' current cpuset, which may be
+  a subset of all online CPUs.
+* All interrupts targeted to this CPU are migrated to a new CPU
+* timers are also migrated to a new CPU
+* Once all services are migrated, kernel calls an arch specific routine
+  ``__cpu_disable()`` to perform arch specific cleanup.
+
+Using the hotplug API
+---------------------
+It is possible to receive notifications once a CPU is offline or onlined. This
+might be important to certain drivers which need to perform some kind of setup
+or clean up functions based on the number of available CPUs: ::
+
+  #include <linux/cpuhotplug.h>
+
+  ret = cpuhp_setup_state(CPUHP_AP_ONLINE_DYN, "X/Y:online",
+                          Y_online, Y_prepare_down);
+
+*X* is the subsystem and *Y* the particular driver. The *Y_online* callback
+will be invoked during registration on all online CPUs. If an error
+occurs during the online callback the *Y_prepare_down* callback will be
+invoked on all CPUs on which the online callback was previously invoked.
+After registration completed, the *Y_online* callback will be invoked
+once a CPU is brought online and *Y_prepare_down* will be invoked when a
+CPU is shutdown. All resources which were previously allocated in
+*Y_online* should be released in *Y_prepare_down*.
+The return value *ret* is negative if an error occurred during the
+registration process. Otherwise a positive value is returned which
+contains the allocated hotplug for dynamically allocated states
+(*CPUHP_AP_ONLINE_DYN*). It will return zero for predefined states.
+
+The callback can be remove by invoking ``cpuhp_remove_state()``. In case of a
+dynamically allocated state (*CPUHP_AP_ONLINE_DYN*) use the returned state.
+During the removal of a hotplug state the teardown callback will be invoked.
+
+Multiple instances
+~~~~~~~~~~~~~~~~~~
+If a driver has multiple instances and each instance needs to perform the
+callback independently then it is likely that a ''multi-state'' should be used.
+First a multi-state state needs to be registered: ::
+
+  ret = cpuhp_setup_state_multi(CPUHP_AP_ONLINE_DYN, "X/Y:online,
+                                Y_online, Y_prepare_down);
+  Y_hp_online = ret;
+
+The ``cpuhp_setup_state_multi()`` behaves similar to ``cpuhp_setup_state()``
+except it prepares the callbacks for a multi state and does not invoke
+the callbacks. This is a one time setup.
+Once a new instance is allocated, you need to register this new instance: ::
+
+  ret = cpuhp_state_add_instance(Y_hp_online, &d->node);
+
+This function will add this instance to your previously allocated
+*Y_hp_online* state and invoke the previously registered callback
+(*Y_online*) on all online CPUs. The *node* element is a ``struct
+hlist_node`` member of your per-instance data structure.
+
+On removal of the instance: ::
+  cpuhp_state_remove_instance(Y_hp_online, &d->node)
+
+should be invoked which will invoke the teardown callback on all online
+CPUs.
+
+Manual setup
+~~~~~~~~~~~~
+Usually it is handy to invoke setup and teardown callbacks on registration or
+removal of a state because usually the operation needs to performed once a CPU
+goes online (offline) and during initial setup (shutdown) of the driver. However
+each registration and removal function is also available with a ``_nocalls``
+suffix which does not invoke the provided callbacks if the invocation of the
+callbacks is not desired. During the manual setup (or teardown) the functions
+``get_online_cpus()`` and ``put_online_cpus()`` should be used to inhibit CPU
+hotplug operations.
+
+
+The ordering of the events
+--------------------------
+The hotplug states are defined in ``include/linux/cpuhotplug.h``:
+
+* The states *CPUHP_OFFLINE* … *CPUHP_AP_OFFLINE* are invoked before the
+  CPU is up.
+* The states *CPUHP_AP_OFFLINE* … *CPUHP_AP_ONLINE* are invoked
+  just the after the CPU has been brought up. The interrupts are off and
+  the scheduler is not yet active on this CPU. Starting with *CPUHP_AP_OFFLINE*
+  the callbacks are invoked on the target CPU.
+* The states between *CPUHP_AP_ONLINE_DYN* and *CPUHP_AP_ONLINE_DYN_END* are
+  reserved for the dynamic allocation.
+* The states are invoked in the reverse order on CPU shutdown starting with
+  *CPUHP_ONLINE* and stopping at *CPUHP_OFFLINE*. Here the callbacks are
+  invoked on the CPU that will be shutdown until *CPUHP_AP_OFFLINE*.
+
+A dynamically allocated state via *CPUHP_AP_ONLINE_DYN* is often enough.
+However if an earlier invocation during the bring up or shutdown is required
+then an explicit state should be acquired. An explicit state might also be
+required if the hotplug event requires specific ordering in respect to
+another hotplug event.
+
+Testing of hotplug states
+=========================
+One way to verify whether a custom state is working as expected or not is to
+shutdown a CPU and then put it online again. It is also possible to put the CPU
+to certain state (for instance *CPUHP_AP_ONLINE*) and then go back to
+*CPUHP_ONLINE*. This would simulate an error one state after *CPUHP_AP_ONLINE*
+which would lead to rollback to the online state.
+
+All registered states are enumerated in ``/sys/devices/system/cpu/hotplug/states``: ::
+
+ $ tail /sys/devices/system/cpu/hotplug/states
+ 138: mm/vmscan:online
+ 139: mm/vmstat:online
+ 140: lib/percpu_cnt:online
+ 141: acpi/cpu-drv:online
+ 142: base/cacheinfo:online
+ 143: virtio/net:online
+ 144: x86/mce:online
+ 145: printk:online
+ 168: sched:active
+ 169: online
+
+To rollback CPU4 to ``lib/percpu_cnt:online`` and back online just issue: ::
+
+  $ cat /sys/devices/system/cpu/cpu4/hotplug/state
+  169
+  $ echo 140 > /sys/devices/system/cpu/cpu4/hotplug/target
+  $ cat /sys/devices/system/cpu/cpu4/hotplug/state
+  140
+
+It is important to note that the teardown callbac of state 140 have been
+invoked. And now get back online: ::
+
+  $ echo 169 > /sys/devices/system/cpu/cpu4/hotplug/target
+  $ cat /sys/devices/system/cpu/cpu4/hotplug/state
+  169
+
+With trace events enabled, the individual steps are visible, too: ::
+
+  #  TASK-PID   CPU#    TIMESTAMP  FUNCTION
+  #     | |       |        |         |
+      bash-394  [001]  22.976: cpuhp_enter: cpu: 0004 target: 140 step: 169 (cpuhp_kick_ap_work)
+   cpuhp/4-31   [004]  22.977: cpuhp_enter: cpu: 0004 target: 140 step: 168 (sched_cpu_deactivate)
+   cpuhp/4-31   [004]  22.990: cpuhp_exit:  cpu: 0004  state: 168 step: 168 ret: 0
+   cpuhp/4-31   [004]  22.991: cpuhp_enter: cpu: 0004 target: 140 step: 144 (mce_cpu_pre_down)
+   cpuhp/4-31   [004]  22.992: cpuhp_exit:  cpu: 0004  state: 144 step: 144 ret: 0
+   cpuhp/4-31   [004]  22.993: cpuhp_multi_enter: cpu: 0004 target: 140 step: 143 (virtnet_cpu_down_prep)
+   cpuhp/4-31   [004]  22.994: cpuhp_exit:  cpu: 0004  state: 143 step: 143 ret: 0
+   cpuhp/4-31   [004]  22.995: cpuhp_enter: cpu: 0004 target: 140 step: 142 (cacheinfo_cpu_pre_down)
+   cpuhp/4-31   [004]  22.996: cpuhp_exit:  cpu: 0004  state: 142 step: 142 ret: 0
+      bash-394  [001]  22.997: cpuhp_exit:  cpu: 0004  state: 140 step: 169 ret: 0
+      bash-394  [005]  95.540: cpuhp_enter: cpu: 0004 target: 169 step: 140 (cpuhp_kick_ap_work)
+   cpuhp/4-31   [004]  95.541: cpuhp_enter: cpu: 0004 target: 169 step: 141 (acpi_soft_cpu_online)
+   cpuhp/4-31   [004]  95.542: cpuhp_exit:  cpu: 0004  state: 141 step: 141 ret: 0
+   cpuhp/4-31   [004]  95.543: cpuhp_enter: cpu: 0004 target: 169 step: 142 (cacheinfo_cpu_online)
+   cpuhp/4-31   [004]  95.544: cpuhp_exit:  cpu: 0004  state: 142 step: 142 ret: 0
+   cpuhp/4-31   [004]  95.545: cpuhp_multi_enter: cpu: 0004 target: 169 step: 143 (virtnet_cpu_online)
+   cpuhp/4-31   [004]  95.546: cpuhp_exit:  cpu: 0004  state: 143 step: 143 ret: 0
+   cpuhp/4-31   [004]  95.547: cpuhp_enter: cpu: 0004 target: 169 step: 144 (mce_cpu_online)
+   cpuhp/4-31   [004]  95.548: cpuhp_exit:  cpu: 0004  state: 144 step: 144 ret: 0
+   cpuhp/4-31   [004]  95.549: cpuhp_enter: cpu: 0004 target: 169 step: 145 (console_cpu_notify)
+   cpuhp/4-31   [004]  95.550: cpuhp_exit:  cpu: 0004  state: 145 step: 145 ret: 0
+   cpuhp/4-31   [004]  95.551: cpuhp_enter: cpu: 0004 target: 169 step: 168 (sched_cpu_activate)
+   cpuhp/4-31   [004]  95.552: cpuhp_exit:  cpu: 0004  state: 168 step: 168 ret: 0
+      bash-394  [005]  95.553: cpuhp_exit:  cpu: 0004  state: 169 step: 140 ret: 0
+
+As it an be seen, CPU4 went down until timestamp 22.996 and then back up until
+95.552. All invoked callbacks including their return codes are visible in the
+trace.
+
+Architecture's requirements
+===========================
+The following functions and configurations are required:
+
+``CONFIG_HOTPLUG_CPU``
+  This entry needs to be enabled in Kconfig
+
+``__cpu_up()``
+  Arch interface to bring up a CPU
+
+``__cpu_disable()``
+  Arch interface to shutdown a CPU, no more interrupts can be handled by the
+  kernel after the routine returns. This includes the shutdown of the timer.
+
+``__cpu_die()``
+  This actually supposed to ensure death of the CPU. Actually look at some
+  example code in other arch that implement CPU hotplug. The processor is taken
+  down from the ``idle()`` loop for that specific architecture. ``__cpu_die()``
+  typically waits for some per_cpu state to be set, to ensure the processor dead
+  routine is called to be sure positively.
+
+User Space Notification
+=======================
+After CPU successfully onlined or offline udev events are sent. A udev rule like: ::
+
+  SUBSYSTEM=="cpu", DRIVERS=="processor", DEVPATH=="/devices/system/cpu/*", RUN+="the_hotplug_receiver.sh"
+
+will receive all events. A script like: ::
+
+  #!/bin/sh
+
+  if [ "${ACTION}" = "offline" ]
+  then
+      echo "CPU ${DEVPATH##*/} offline"
+
+  elif [ "${ACTION}" = "online" ]
+  then
+      echo "CPU ${DEVPATH##*/} online"
+
+  fi
+
+can process the event further.
+
+Kernel Inline Documentations Reference
+======================================
+
+.. kernel-doc:: include/linux/cpuhotplug.h
diff --git a/Documentation/core-api/debug-objects.rst b/Documentation/core-api/debug-objects.rst
new file mode 100644
index 000000000..ac926fd55
--- /dev/null
+++ b/Documentation/core-api/debug-objects.rst
@@ -0,0 +1,310 @@
+============================================
+The object-lifetime debugging infrastructure
+============================================
+
+:Author: Thomas Gleixner
+
+Introduction
+============
+
+debugobjects is a generic infrastructure to track the life time of
+kernel objects and validate the operations on those.
+
+debugobjects is useful to check for the following error patterns:
+
+-  Activation of uninitialized objects
+
+-  Initialization of active objects
+
+-  Usage of freed/destroyed objects
+
+debugobjects is not changing the data structure of the real object so it
+can be compiled in with a minimal runtime impact and enabled on demand
+with a kernel command line option.
+
+Howto use debugobjects
+======================
+
+A kernel subsystem needs to provide a data structure which describes the
+object type and add calls into the debug code at appropriate places. The
+data structure to describe the object type needs at minimum the name of
+the object type. Optional functions can and should be provided to fixup
+detected problems so the kernel can continue to work and the debug
+information can be retrieved from a live system instead of hard core
+debugging with serial consoles and stack trace transcripts from the
+monitor.
+
+The debug calls provided by debugobjects are:
+
+-  debug_object_init
+
+-  debug_object_init_on_stack
+
+-  debug_object_activate
+
+-  debug_object_deactivate
+
+-  debug_object_destroy
+
+-  debug_object_free
+
+-  debug_object_assert_init
+
+Each of these functions takes the address of the real object and a
+pointer to the object type specific debug description structure.
+
+Each detected error is reported in the statistics and a limited number
+of errors are printk'ed including a full stack trace.
+
+The statistics are available via /sys/kernel/debug/debug_objects/stats.
+They provide information about the number of warnings and the number of
+successful fixups along with information about the usage of the internal
+tracking objects and the state of the internal tracking objects pool.
+
+Debug functions
+===============
+
+.. kernel-doc:: lib/debugobjects.c
+   :functions: debug_object_init
+
+This function is called whenever the initialization function of a real
+object is called.
+
+When the real object is already tracked by debugobjects it is checked,
+whether the object can be initialized. Initializing is not allowed for
+active and destroyed objects. When debugobjects detects an error, then
+it calls the fixup_init function of the object type description
+structure if provided by the caller. The fixup function can correct the
+problem before the real initialization of the object happens. E.g. it
+can deactivate an active object in order to prevent damage to the
+subsystem.
+
+When the real object is not yet tracked by debugobjects, debugobjects
+allocates a tracker object for the real object and sets the tracker
+object state to ODEBUG_STATE_INIT. It verifies that the object is not
+on the callers stack. If it is on the callers stack then a limited
+number of warnings including a full stack trace is printk'ed. The
+calling code must use debug_object_init_on_stack() and remove the
+object before leaving the function which allocated it. See next section.
+
+.. kernel-doc:: lib/debugobjects.c
+   :functions: debug_object_init_on_stack
+
+This function is called whenever the initialization function of a real
+object which resides on the stack is called.
+
+When the real object is already tracked by debugobjects it is checked,
+whether the object can be initialized. Initializing is not allowed for
+active and destroyed objects. When debugobjects detects an error, then
+it calls the fixup_init function of the object type description
+structure if provided by the caller. The fixup function can correct the
+problem before the real initialization of the object happens. E.g. it
+can deactivate an active object in order to prevent damage to the
+subsystem.
+
+When the real object is not yet tracked by debugobjects debugobjects
+allocates a tracker object for the real object and sets the tracker
+object state to ODEBUG_STATE_INIT. It verifies that the object is on
+the callers stack.
+
+An object which is on the stack must be removed from the tracker by
+calling debug_object_free() before the function which allocates the
+object returns. Otherwise we keep track of stale objects.
+
+.. kernel-doc:: lib/debugobjects.c
+   :functions: debug_object_activate
+
+This function is called whenever the activation function of a real
+object is called.
+
+When the real object is already tracked by debugobjects it is checked,
+whether the object can be activated. Activating is not allowed for
+active and destroyed objects. When debugobjects detects an error, then
+it calls the fixup_activate function of the object type description
+structure if provided by the caller. The fixup function can correct the
+problem before the real activation of the object happens. E.g. it can
+deactivate an active object in order to prevent damage to the subsystem.
+
+When the real object is not yet tracked by debugobjects then the
+fixup_activate function is called if available. This is necessary to
+allow the legitimate activation of statically allocated and initialized
+objects. The fixup function checks whether the object is valid and calls
+the debug_objects_init() function to initialize the tracking of this
+object.
+
+When the activation is legitimate, then the state of the associated
+tracker object is set to ODEBUG_STATE_ACTIVE.
+
+
+.. kernel-doc:: lib/debugobjects.c
+   :functions: debug_object_deactivate
+
+This function is called whenever the deactivation function of a real
+object is called.
+
+When the real object is tracked by debugobjects it is checked, whether
+the object can be deactivated. Deactivating is not allowed for untracked
+or destroyed objects.
+
+When the deactivation is legitimate, then the state of the associated
+tracker object is set to ODEBUG_STATE_INACTIVE.
+
+.. kernel-doc:: lib/debugobjects.c
+   :functions: debug_object_destroy
+
+This function is called to mark an object destroyed. This is useful to
+prevent the usage of invalid objects, which are still available in
+memory: either statically allocated objects or objects which are freed
+later.
+
+When the real object is tracked by debugobjects it is checked, whether
+the object can be destroyed. Destruction is not allowed for active and
+destroyed objects. When debugobjects detects an error, then it calls the
+fixup_destroy function of the object type description structure if
+provided by the caller. The fixup function can correct the problem
+before the real destruction of the object happens. E.g. it can
+deactivate an active object in order to prevent damage to the subsystem.
+
+When the destruction is legitimate, then the state of the associated
+tracker object is set to ODEBUG_STATE_DESTROYED.
+
+.. kernel-doc:: lib/debugobjects.c
+   :functions: debug_object_free
+
+This function is called before an object is freed.
+
+When the real object is tracked by debugobjects it is checked, whether
+the object can be freed. Free is not allowed for active objects. When
+debugobjects detects an error, then it calls the fixup_free function of
+the object type description structure if provided by the caller. The
+fixup function can correct the problem before the real free of the
+object happens. E.g. it can deactivate an active object in order to
+prevent damage to the subsystem.
+
+Note that debug_object_free removes the object from the tracker. Later
+usage of the object is detected by the other debug checks.
+
+
+.. kernel-doc:: lib/debugobjects.c
+   :functions: debug_object_assert_init
+
+This function is called to assert that an object has been initialized.
+
+When the real object is not tracked by debugobjects, it calls
+fixup_assert_init of the object type description structure provided by
+the caller, with the hardcoded object state ODEBUG_NOT_AVAILABLE. The
+fixup function can correct the problem by calling debug_object_init
+and other specific initializing functions.
+
+When the real object is already tracked by debugobjects it is ignored.
+
+Fixup functions
+===============
+
+Debug object type description structure
+---------------------------------------
+
+.. kernel-doc:: include/linux/debugobjects.h
+   :internal:
+
+fixup_init
+-----------
+
+This function is called from the debug code whenever a problem in
+debug_object_init is detected. The function takes the address of the
+object and the state which is currently recorded in the tracker.
+
+Called from debug_object_init when the object state is:
+
+-  ODEBUG_STATE_ACTIVE
+
+The function returns true when the fixup was successful, otherwise
+false. The return value is used to update the statistics.
+
+Note, that the function needs to call the debug_object_init() function
+again, after the damage has been repaired in order to keep the state
+consistent.
+
+fixup_activate
+---------------
+
+This function is called from the debug code whenever a problem in
+debug_object_activate is detected.
+
+Called from debug_object_activate when the object state is:
+
+-  ODEBUG_STATE_NOTAVAILABLE
+
+-  ODEBUG_STATE_ACTIVE
+
+The function returns true when the fixup was successful, otherwise
+false. The return value is used to update the statistics.
+
+Note that the function needs to call the debug_object_activate()
+function again after the damage has been repaired in order to keep the
+state consistent.
+
+The activation of statically initialized objects is a special case. When
+debug_object_activate() has no tracked object for this object address
+then fixup_activate() is called with object state
+ODEBUG_STATE_NOTAVAILABLE. The fixup function needs to check whether
+this is a legitimate case of a statically initialized object or not. In
+case it is it calls debug_object_init() and debug_object_activate()
+to make the object known to the tracker and marked active. In this case
+the function should return false because this is not a real fixup.
+
+fixup_destroy
+--------------
+
+This function is called from the debug code whenever a problem in
+debug_object_destroy is detected.
+
+Called from debug_object_destroy when the object state is:
+
+-  ODEBUG_STATE_ACTIVE
+
+The function returns true when the fixup was successful, otherwise
+false. The return value is used to update the statistics.
+
+fixup_free
+-----------
+
+This function is called from the debug code whenever a problem in
+debug_object_free is detected. Further it can be called from the debug
+checks in kfree/vfree, when an active object is detected from the
+debug_check_no_obj_freed() sanity checks.
+
+Called from debug_object_free() or debug_check_no_obj_freed() when
+the object state is:
+
+-  ODEBUG_STATE_ACTIVE
+
+The function returns true when the fixup was successful, otherwise
+false. The return value is used to update the statistics.
+
+fixup_assert_init
+-------------------
+
+This function is called from the debug code whenever a problem in
+debug_object_assert_init is detected.
+
+Called from debug_object_assert_init() with a hardcoded state
+ODEBUG_STATE_NOTAVAILABLE when the object is not found in the debug
+bucket.
+
+The function returns true when the fixup was successful, otherwise
+false. The return value is used to update the statistics.
+
+Note, this function should make sure debug_object_init() is called
+before returning.
+
+The handling of statically initialized objects is a special case. The
+fixup function should check if this is a legitimate case of a statically
+initialized object or not. In this case only debug_object_init()
+should be called to make the object known to the tracker. Then the
+function should return false because this is not a real fixup.
+
+Known Bugs And Assumptions
+==========================
+
+None (knock on wood).
diff --git a/Documentation/core-api/debugging-via-ohci1394.rst b/Documentation/core-api/debugging-via-ohci1394.rst
new file mode 100644
index 000000000..981ad4f89
--- /dev/null
+++ b/Documentation/core-api/debugging-via-ohci1394.rst
@@ -0,0 +1,185 @@
+===========================================================================
+Using physical DMA provided by OHCI-1394 FireWire controllers for debugging
+===========================================================================
+
+Introduction
+------------
+
+Basically all FireWire controllers which are in use today are compliant
+to the OHCI-1394 specification which defines the controller to be a PCI
+bus master which uses DMA to offload data transfers from the CPU and has
+a "Physical Response Unit" which executes specific requests by employing
+PCI-Bus master DMA after applying filters defined by the OHCI-1394 driver.
+
+Once properly configured, remote machines can send these requests to
+ask the OHCI-1394 controller to perform read and write requests on
+physical system memory and, for read requests, send the result of
+the physical memory read back to the requester.
+
+With that, it is possible to debug issues by reading interesting memory
+locations such as buffers like the printk buffer or the process table.
+
+Retrieving a full system memory dump is also possible over the FireWire,
+using data transfer rates in the order of 10MB/s or more.
+
+With most FireWire controllers, memory access is limited to the low 4 GB
+of physical address space.  This can be a problem on IA64 machines where
+memory is located mostly above that limit, but it is rarely a problem on
+more common hardware such as x86, x86-64 and PowerPC.
+
+At least LSI FW643e and FW643e2 controllers are known to support access to
+physical addresses above 4 GB, but this feature is currently not enabled by
+Linux.
+
+Together with a early initialization of the OHCI-1394 controller for debugging,
+this facility proved most useful for examining long debugs logs in the printk
+buffer on to debug early boot problems in areas like ACPI where the system
+fails to boot and other means for debugging (serial port) are either not
+available (notebooks) or too slow for extensive debug information (like ACPI).
+
+Drivers
+-------
+
+The firewire-ohci driver in drivers/firewire uses filtered physical
+DMA by default, which is more secure but not suitable for remote debugging.
+Pass the remote_dma=1 parameter to the driver to get unfiltered physical DMA.
+
+Because the firewire-ohci driver depends on the PCI enumeration to be
+completed, an initialization routine which runs pretty early has been
+implemented for x86.  This routine runs long before console_init() can be
+called, i.e. before the printk buffer appears on the console.
+
+To activate it, enable CONFIG_PROVIDE_OHCI1394_DMA_INIT (Kernel hacking menu:
+Remote debugging over FireWire early on boot) and pass the parameter
+"ohci1394_dma=early" to the recompiled kernel on boot.
+
+Tools
+-----
+
+firescope - Originally developed by Benjamin Herrenschmidt, Andi Kleen ported
+it from PowerPC to x86 and x86_64 and added functionality, firescope can now
+be used to view the printk buffer of a remote machine, even with live update.
+
+Bernhard Kaindl enhanced firescope to support accessing 64-bit machines
+from 32-bit firescope and vice versa:
+- http://v3.sk/~lkundrak/firescope/
+
+and he implemented fast system dump (alpha version - read README.txt):
+- http://halobates.de/firewire/firedump-0.1.tar.bz2
+
+There is also a gdb proxy for firewire which allows to use gdb to access
+data which can be referenced from symbols found by gdb in vmlinux:
+- http://halobates.de/firewire/fireproxy-0.33.tar.bz2
+
+The latest version of this gdb proxy (fireproxy-0.34) can communicate (not
+yet stable) with kgdb over an memory-based communication module (kgdbom).
+
+Getting Started
+---------------
+
+The OHCI-1394 specification regulates that the OHCI-1394 controller must
+disable all physical DMA on each bus reset.
+
+This means that if you want to debug an issue in a system state where
+interrupts are disabled and where no polling of the OHCI-1394 controller
+for bus resets takes place, you have to establish any FireWire cable
+connections and fully initialize all FireWire hardware __before__ the
+system enters such state.
+
+Step-by-step instructions for using firescope with early OHCI initialization:
+
+1) Verify that your hardware is supported:
+
+   Load the firewire-ohci module and check your kernel logs.
+   You should see a line similar to::
+
+     firewire_ohci 0000:15:00.1: added OHCI v1.0 device as card 2, 4 IR + 4 IT
+     ... contexts, quirks 0x11
+
+   when loading the driver. If you have no supported controller, many PCI,
+   CardBus and even some Express cards which are fully compliant to OHCI-1394
+   specification are available. If it requires no driver for Windows operating
+   systems, it most likely is. Only specialized shops have cards which are not
+   compliant, they are based on TI PCILynx chips and require drivers for Windows
+   operating systems.
+
+   The mentioned kernel log message contains the string "physUB" if the
+   controller implements a writable Physical Upper Bound register.  This is
+   required for physical DMA above 4 GB (but not utilized by Linux yet).
+
+2) Establish a working FireWire cable connection:
+
+   Any FireWire cable, as long at it provides electrically and mechanically
+   stable connection and has matching connectors (there are small 4-pin and
+   large 6-pin FireWire ports) will do.
+
+   If an driver is running on both machines you should see a line like::
+
+     firewire_core 0000:15:00.1: created device fw1: GUID 00061b0020105917, S400
+
+   on both machines in the kernel log when the cable is plugged in
+   and connects the two machines.
+
+3) Test physical DMA using firescope:
+
+   On the debug host, make sure that /dev/fw* is accessible,
+   then start firescope::
+
+	$ firescope
+	Port 0 (/dev/fw1) opened, 2 nodes detected
+
+	FireScope
+	---------
+	Target : <unspecified>
+	Gen    : 1
+	[Ctrl-T] choose target
+	[Ctrl-H] this menu
+	[Ctrl-Q] quit
+
+    ------> Press Ctrl-T now, the output should be similar to:
+
+	2 nodes available, local node is: 0
+	 0: ffc0, uuid: 00000000 00000000 [LOCAL]
+	 1: ffc1, uuid: 00279000 ba4bb801
+
+   Besides the [LOCAL] node, it must show another node without error message.
+
+4) Prepare for debugging with early OHCI-1394 initialization:
+
+   4.1) Kernel compilation and installation on debug target
+
+   Compile the kernel to be debugged with CONFIG_PROVIDE_OHCI1394_DMA_INIT
+   (Kernel hacking: Provide code for enabling DMA over FireWire early on boot)
+   enabled and install it on the machine to be debugged (debug target).
+
+   4.2) Transfer the System.map of the debugged kernel to the debug host
+
+   Copy the System.map of the kernel be debugged to the debug host (the host
+   which is connected to the debugged machine over the FireWire cable).
+
+5) Retrieving the printk buffer contents:
+
+   With the FireWire cable connected, the OHCI-1394 driver on the debugging
+   host loaded, reboot the debugged machine, booting the kernel which has
+   CONFIG_PROVIDE_OHCI1394_DMA_INIT enabled, with the option ohci1394_dma=early.
+
+   Then, on the debugging host, run firescope, for example by using -A::
+
+	firescope -A System.map-of-debug-target-kernel
+
+   Note: -A automatically attaches to the first non-local node. It only works
+   reliably if only connected two machines are connected using FireWire.
+
+   After having attached to the debug target, press Ctrl-D to view the
+   complete printk buffer or Ctrl-U to enter auto update mode and get an
+   updated live view of recent kernel messages logged on the debug target.
+
+   Call "firescope -h" to get more information on firescope's options.
+
+Notes
+-----
+
+Documentation and specifications: http://halobates.de/firewire/
+
+FireWire is a trademark of Apple Inc. - for more information please refer to:
+https://en.wikipedia.org/wiki/FireWire
diff --git a/Documentation/core-api/dma-api-howto.rst b/Documentation/core-api/dma-api-howto.rst
new file mode 100644
index 000000000..358d49545
--- /dev/null
+++ b/Documentation/core-api/dma-api-howto.rst
@@ -0,0 +1,929 @@
+=========================
+Dynamic DMA mapping Guide
+=========================
+
+:Author: David S. Miller <davem@redhat.com>
+:Author: Richard Henderson <rth@cygnus.com>
+:Author: Jakub Jelinek <jakub@redhat.com>
+
+This is a guide to device driver writers on how to use the DMA API
+with example pseudo-code.  For a concise description of the API, see
+DMA-API.txt.
+
+CPU and DMA addresses
+=====================
+
+There are several kinds of addresses involved in the DMA API, and it's
+important to understand the differences.
+
+The kernel normally uses virtual addresses.  Any address returned by
+kmalloc(), vmalloc(), and similar interfaces is a virtual address and can
+be stored in a ``void *``.
+
+The virtual memory system (TLB, page tables, etc.) translates virtual
+addresses to CPU physical addresses, which are stored as "phys_addr_t" or
+"resource_size_t".  The kernel manages device resources like registers as
+physical addresses.  These are the addresses in /proc/iomem.  The physical
+address is not directly useful to a driver; it must use ioremap() to map
+the space and produce a virtual address.
+
+I/O devices use a third kind of address: a "bus address".  If a device has
+registers at an MMIO address, or if it performs DMA to read or write system
+memory, the addresses used by the device are bus addresses.  In some
+systems, bus addresses are identical to CPU physical addresses, but in
+general they are not.  IOMMUs and host bridges can produce arbitrary
+mappings between physical and bus addresses.
+
+From a device's point of view, DMA uses the bus address space, but it may
+be restricted to a subset of that space.  For example, even if a system
+supports 64-bit addresses for main memory and PCI BARs, it may use an IOMMU
+so devices only need to use 32-bit DMA addresses.
+
+Here's a picture and some examples::
+
+               CPU                  CPU                  Bus
+             Virtual              Physical             Address
+             Address              Address               Space
+              Space                Space
+
+            +-------+             +------+             +------+
+            |       |             |MMIO  |   Offset    |      |
+            |       |  Virtual    |Space |   applied   |      |
+          C +-------+ --------> B +------+ ----------> +------+ A
+            |       |  mapping    |      |   by host   |      |
+  +-----+   |       |             |      |   bridge    |      |   +--------+
+  |     |   |       |             +------+             |      |   |        |
+  | CPU |   |       |             | RAM  |             |      |   | Device |
+  |     |   |       |             |      |             |      |   |        |
+  +-----+   +-------+             +------+             +------+   +--------+
+            |       |  Virtual    |Buffer|   Mapping   |      |
+          X +-------+ --------> Y +------+ <---------- +------+ Z
+            |       |  mapping    | RAM  |   by IOMMU
+            |       |             |      |
+            |       |             |      |
+            +-------+             +------+
+
+During the enumeration process, the kernel learns about I/O devices and
+their MMIO space and the host bridges that connect them to the system.  For
+example, if a PCI device has a BAR, the kernel reads the bus address (A)
+from the BAR and converts it to a CPU physical address (B).  The address B
+is stored in a struct resource and usually exposed via /proc/iomem.  When a
+driver claims a device, it typically uses ioremap() to map physical address
+B at a virtual address (C).  It can then use, e.g., ioread32(C), to access
+the device registers at bus address A.
+
+If the device supports DMA, the driver sets up a buffer using kmalloc() or
+a similar interface, which returns a virtual address (X).  The virtual
+memory system maps X to a physical address (Y) in system RAM.  The driver
+can use virtual address X to access the buffer, but the device itself
+cannot because DMA doesn't go through the CPU virtual memory system.
+
+In some simple systems, the device can do DMA directly to physical address
+Y.  But in many others, there is IOMMU hardware that translates DMA
+addresses to physical addresses, e.g., it translates Z to Y.  This is part
+of the reason for the DMA API: the driver can give a virtual address X to
+an interface like dma_map_single(), which sets up any required IOMMU
+mapping and returns the DMA address Z.  The driver then tells the device to
+do DMA to Z, and the IOMMU maps it to the buffer at address Y in system
+RAM.
+
+So that Linux can use the dynamic DMA mapping, it needs some help from the
+drivers, namely it has to take into account that DMA addresses should be
+mapped only for the time they are actually used and unmapped after the DMA
+transfer.
+
+The following API will work of course even on platforms where no such
+hardware exists.
+
+Note that the DMA API works with any bus independent of the underlying
+microprocessor architecture. You should use the DMA API rather than the
+bus-specific DMA API, i.e., use the dma_map_*() interfaces rather than the
+pci_map_*() interfaces.
+
+First of all, you should make sure::
+
+	#include <linux/dma-mapping.h>
+
+is in your driver, which provides the definition of dma_addr_t.  This type
+can hold any valid DMA address for the platform and should be used
+everywhere you hold a DMA address returned from the DMA mapping functions.
+
+What memory is DMA'able?
+========================
+
+The first piece of information you must know is what kernel memory can
+be used with the DMA mapping facilities.  There has been an unwritten
+set of rules regarding this, and this text is an attempt to finally
+write them down.
+
+If you acquired your memory via the page allocator
+(i.e. __get_free_page*()) or the generic memory allocators
+(i.e. kmalloc() or kmem_cache_alloc()) then you may DMA to/from
+that memory using the addresses returned from those routines.
+
+This means specifically that you may _not_ use the memory/addresses
+returned from vmalloc() for DMA.  It is possible to DMA to the
+_underlying_ memory mapped into a vmalloc() area, but this requires
+walking page tables to get the physical addresses, and then
+translating each of those pages back to a kernel address using
+something like __va().  [ EDIT: Update this when we integrate
+Gerd Knorr's generic code which does this. ]
+
+This rule also means that you may use neither kernel image addresses
+(items in data/text/bss segments), nor module image addresses, nor
+stack addresses for DMA.  These could all be mapped somewhere entirely
+different than the rest of physical memory.  Even if those classes of
+memory could physically work with DMA, you'd need to ensure the I/O
+buffers were cacheline-aligned.  Without that, you'd see cacheline
+sharing problems (data corruption) on CPUs with DMA-incoherent caches.
+(The CPU could write to one word, DMA would write to a different one
+in the same cache line, and one of them could be overwritten.)
+
+Also, this means that you cannot take the return of a kmap()
+call and DMA to/from that.  This is similar to vmalloc().
+
+What about block I/O and networking buffers?  The block I/O and
+networking subsystems make sure that the buffers they use are valid
+for you to DMA from/to.
+
+DMA addressing capabilities
+===========================
+
+By default, the kernel assumes that your device can address 32-bits of DMA
+addressing.  For a 64-bit capable device, this needs to be increased, and for
+a device with limitations, it needs to be decreased.
+
+Special note about PCI: PCI-X specification requires PCI-X devices to support
+64-bit addressing (DAC) for all transactions.  And at least one platform (SGI
+SN2) requires 64-bit consistent allocations to operate correctly when the IO
+bus is in PCI-X mode.
+
+For correct operation, you must set the DMA mask to inform the kernel about
+your devices DMA addressing capabilities.
+
+This is performed via a call to dma_set_mask_and_coherent()::
+
+	int dma_set_mask_and_coherent(struct device *dev, u64 mask);
+
+which will set the mask for both streaming and coherent APIs together.  If you
+have some special requirements, then the following two separate calls can be
+used instead:
+
+	The setup for streaming mappings is performed via a call to
+	dma_set_mask()::
+
+		int dma_set_mask(struct device *dev, u64 mask);
+
+	The setup for consistent allocations is performed via a call
+	to dma_set_coherent_mask()::
+
+		int dma_set_coherent_mask(struct device *dev, u64 mask);
+
+Here, dev is a pointer to the device struct of your device, and mask is a bit
+mask describing which bits of an address your device supports.  Often the
+device struct of your device is embedded in the bus-specific device struct of
+your device.  For example, &pdev->dev is a pointer to the device struct of a
+PCI device (pdev is a pointer to the PCI device struct of your device).
+
+These calls usually return zero to indicated your device can perform DMA
+properly on the machine given the address mask you provided, but they might
+return an error if the mask is too small to be supportable on the given
+system.  If it returns non-zero, your device cannot perform DMA properly on
+this platform, and attempting to do so will result in undefined behavior.
+You must not use DMA on this device unless the dma_set_mask family of
+functions has returned success.
+
+This means that in the failure case, you have two options:
+
+1) Use some non-DMA mode for data transfer, if possible.
+2) Ignore this device and do not initialize it.
+
+It is recommended that your driver print a kernel KERN_WARNING message when
+setting the DMA mask fails.  In this manner, if a user of your driver reports
+that performance is bad or that the device is not even detected, you can ask
+them for the kernel messages to find out exactly why.
+
+The standard 64-bit addressing device would do something like this::
+
+	if (dma_set_mask_and_coherent(dev, DMA_BIT_MASK(64))) {
+		dev_warn(dev, "mydev: No suitable DMA available\n");
+		goto ignore_this_device;
+	}
+
+If the device only supports 32-bit addressing for descriptors in the
+coherent allocations, but supports full 64-bits for streaming mappings
+it would look like this::
+
+	if (dma_set_mask(dev, DMA_BIT_MASK(64))) {
+		dev_warn(dev, "mydev: No suitable DMA available\n");
+		goto ignore_this_device;
+	}
+
+The coherent mask will always be able to set the same or a smaller mask as
+the streaming mask. However for the rare case that a device driver only
+uses consistent allocations, one would have to check the return value from
+dma_set_coherent_mask().
+
+Finally, if your device can only drive the low 24-bits of
+address you might do something like::
+
+	if (dma_set_mask(dev, DMA_BIT_MASK(24))) {
+		dev_warn(dev, "mydev: 24-bit DMA addressing not available\n");
+		goto ignore_this_device;
+	}
+
+When dma_set_mask() or dma_set_mask_and_coherent() is successful, and
+returns zero, the kernel saves away this mask you have provided.  The
+kernel will use this information later when you make DMA mappings.
+
+There is a case which we are aware of at this time, which is worth
+mentioning in this documentation.  If your device supports multiple
+functions (for example a sound card provides playback and record
+functions) and the various different functions have _different_
+DMA addressing limitations, you may wish to probe each mask and
+only provide the functionality which the machine can handle.  It
+is important that the last call to dma_set_mask() be for the
+most specific mask.
+
+Here is pseudo-code showing how this might be done::
+
+	#define PLAYBACK_ADDRESS_BITS	DMA_BIT_MASK(32)
+	#define RECORD_ADDRESS_BITS	DMA_BIT_MASK(24)
+
+	struct my_sound_card *card;
+	struct device *dev;
+
+	...
+	if (!dma_set_mask(dev, PLAYBACK_ADDRESS_BITS)) {
+		card->playback_enabled = 1;
+	} else {
+		card->playback_enabled = 0;
+		dev_warn(dev, "%s: Playback disabled due to DMA limitations\n",
+		       card->name);
+	}
+	if (!dma_set_mask(dev, RECORD_ADDRESS_BITS)) {
+		card->record_enabled = 1;
+	} else {
+		card->record_enabled = 0;
+		dev_warn(dev, "%s: Record disabled due to DMA limitations\n",
+		       card->name);
+	}
+
+A sound card was used as an example here because this genre of PCI
+devices seems to be littered with ISA chips given a PCI front end,
+and thus retaining the 16MB DMA addressing limitations of ISA.
+
+Types of DMA mappings
+=====================
+
+There are two types of DMA mappings:
+
+- Consistent DMA mappings which are usually mapped at driver
+  initialization, unmapped at the end and for which the hardware should
+  guarantee that the device and the CPU can access the data
+  in parallel and will see updates made by each other without any
+  explicit software flushing.
+
+  Think of "consistent" as "synchronous" or "coherent".
+
+  The current default is to return consistent memory in the low 32
+  bits of the DMA space.  However, for future compatibility you should
+  set the consistent mask even if this default is fine for your
+  driver.
+
+  Good examples of what to use consistent mappings for are:
+
+	- Network card DMA ring descriptors.
+	- SCSI adapter mailbox command data structures.
+	- Device firmware microcode executed out of
+	  main memory.
+
+  The invariant these examples all require is that any CPU store
+  to memory is immediately visible to the device, and vice
+  versa.  Consistent mappings guarantee this.
+
+  .. important::
+
+	     Consistent DMA memory does not preclude the usage of
+	     proper memory barriers.  The CPU may reorder stores to
+	     consistent memory just as it may normal memory.  Example:
+	     if it is important for the device to see the first word
+	     of a descriptor updated before the second, you must do
+	     something like::
+
+		desc->word0 = address;
+		wmb();
+		desc->word1 = DESC_VALID;
+
+             in order to get correct behavior on all platforms.
+
+	     Also, on some platforms your driver may need to flush CPU write
+	     buffers in much the same way as it needs to flush write buffers
+	     found in PCI bridges (such as by reading a register's value
+	     after writing it).
+
+- Streaming DMA mappings which are usually mapped for one DMA
+  transfer, unmapped right after it (unless you use dma_sync_* below)
+  and for which hardware can optimize for sequential accesses.
+
+  Think of "streaming" as "asynchronous" or "outside the coherency
+  domain".
+
+  Good examples of what to use streaming mappings for are:
+
+	- Networking buffers transmitted/received by a device.
+	- Filesystem buffers written/read by a SCSI device.
+
+  The interfaces for using this type of mapping were designed in
+  such a way that an implementation can make whatever performance
+  optimizations the hardware allows.  To this end, when using
+  such mappings you must be explicit about what you want to happen.
+
+Neither type of DMA mapping has alignment restrictions that come from
+the underlying bus, although some devices may have such restrictions.
+Also, systems with caches that aren't DMA-coherent will work better
+when the underlying buffers don't share cache lines with other data.
+
+
+Using Consistent DMA mappings
+=============================
+
+To allocate and map large (PAGE_SIZE or so) consistent DMA regions,
+you should do::
+
+	dma_addr_t dma_handle;
+
+	cpu_addr = dma_alloc_coherent(dev, size, &dma_handle, gfp);
+
+where device is a ``struct device *``. This may be called in interrupt
+context with the GFP_ATOMIC flag.
+
+Size is the length of the region you want to allocate, in bytes.
+
+This routine will allocate RAM for that region, so it acts similarly to
+__get_free_pages() (but takes size instead of a page order).  If your
+driver needs regions sized smaller than a page, you may prefer using
+the dma_pool interface, described below.
+
+The consistent DMA mapping interfaces, will by default return a DMA address
+which is 32-bit addressable.  Even if the device indicates (via the DMA mask)
+that it may address the upper 32-bits, consistent allocation will only
+return > 32-bit addresses for DMA if the consistent DMA mask has been
+explicitly changed via dma_set_coherent_mask().  This is true of the
+dma_pool interface as well.
+
+dma_alloc_coherent() returns two values: the virtual address which you
+can use to access it from the CPU and dma_handle which you pass to the
+card.
+
+The CPU virtual address and the DMA address are both
+guaranteed to be aligned to the smallest PAGE_SIZE order which
+is greater than or equal to the requested size.  This invariant
+exists (for example) to guarantee that if you allocate a chunk
+which is smaller than or equal to 64 kilobytes, the extent of the
+buffer you receive will not cross a 64K boundary.
+
+To unmap and free such a DMA region, you call::
+
+	dma_free_coherent(dev, size, cpu_addr, dma_handle);
+
+where dev, size are the same as in the above call and cpu_addr and
+dma_handle are the values dma_alloc_coherent() returned to you.
+This function may not be called in interrupt context.
+
+If your driver needs lots of smaller memory regions, you can write
+custom code to subdivide pages returned by dma_alloc_coherent(),
+or you can use the dma_pool API to do that.  A dma_pool is like
+a kmem_cache, but it uses dma_alloc_coherent(), not __get_free_pages().
+Also, it understands common hardware constraints for alignment,
+like queue heads needing to be aligned on N byte boundaries.
+
+Create a dma_pool like this::
+
+	struct dma_pool *pool;
+
+	pool = dma_pool_create(name, dev, size, align, boundary);
+
+The "name" is for diagnostics (like a kmem_cache name); dev and size
+are as above.  The device's hardware alignment requirement for this
+type of data is "align" (which is expressed in bytes, and must be a
+power of two).  If your device has no boundary crossing restrictions,
+pass 0 for boundary; passing 4096 says memory allocated from this pool
+must not cross 4KByte boundaries (but at that time it may be better to
+use dma_alloc_coherent() directly instead).
+
+Allocate memory from a DMA pool like this::
+
+	cpu_addr = dma_pool_alloc(pool, flags, &dma_handle);
+
+flags are GFP_KERNEL if blocking is permitted (not in_interrupt nor
+holding SMP locks), GFP_ATOMIC otherwise.  Like dma_alloc_coherent(),
+this returns two values, cpu_addr and dma_handle.
+
+Free memory that was allocated from a dma_pool like this::
+
+	dma_pool_free(pool, cpu_addr, dma_handle);
+
+where pool is what you passed to dma_pool_alloc(), and cpu_addr and
+dma_handle are the values dma_pool_alloc() returned. This function
+may be called in interrupt context.
+
+Destroy a dma_pool by calling::
+
+	dma_pool_destroy(pool);
+
+Make sure you've called dma_pool_free() for all memory allocated
+from a pool before you destroy the pool. This function may not
+be called in interrupt context.
+
+DMA Direction
+=============
+
+The interfaces described in subsequent portions of this document
+take a DMA direction argument, which is an integer and takes on
+one of the following values::
+
+ DMA_BIDIRECTIONAL
+ DMA_TO_DEVICE
+ DMA_FROM_DEVICE
+ DMA_NONE
+
+You should provide the exact DMA direction if you know it.
+
+DMA_TO_DEVICE means "from main memory to the device"
+DMA_FROM_DEVICE means "from the device to main memory"
+It is the direction in which the data moves during the DMA
+transfer.
+
+You are _strongly_ encouraged to specify this as precisely
+as you possibly can.
+
+If you absolutely cannot know the direction of the DMA transfer,
+specify DMA_BIDIRECTIONAL.  It means that the DMA can go in
+either direction.  The platform guarantees that you may legally
+specify this, and that it will work, but this may be at the
+cost of performance for example.
+
+The value DMA_NONE is to be used for debugging.  One can
+hold this in a data structure before you come to know the
+precise direction, and this will help catch cases where your
+direction tracking logic has failed to set things up properly.
+
+Another advantage of specifying this value precisely (outside of
+potential platform-specific optimizations of such) is for debugging.
+Some platforms actually have a write permission boolean which DMA
+mappings can be marked with, much like page protections in the user
+program address space.  Such platforms can and do report errors in the
+kernel logs when the DMA controller hardware detects violation of the
+permission setting.
+
+Only streaming mappings specify a direction, consistent mappings
+implicitly have a direction attribute setting of
+DMA_BIDIRECTIONAL.
+
+The SCSI subsystem tells you the direction to use in the
+'sc_data_direction' member of the SCSI command your driver is
+working on.
+
+For Networking drivers, it's a rather simple affair.  For transmit
+packets, map/unmap them with the DMA_TO_DEVICE direction
+specifier.  For receive packets, just the opposite, map/unmap them
+with the DMA_FROM_DEVICE direction specifier.
+
+Using Streaming DMA mappings
+============================
+
+The streaming DMA mapping routines can be called from interrupt
+context.  There are two versions of each map/unmap, one which will
+map/unmap a single memory region, and one which will map/unmap a
+scatterlist.
+
+To map a single region, you do::
+
+	struct device *dev = &my_dev->dev;
+	dma_addr_t dma_handle;
+	void *addr = buffer->ptr;
+	size_t size = buffer->len;
+
+	dma_handle = dma_map_single(dev, addr, size, direction);
+	if (dma_mapping_error(dev, dma_handle)) {
+		/*
+		 * reduce current DMA mapping usage,
+		 * delay and try again later or
+		 * reset driver.
+		 */
+		goto map_error_handling;
+	}
+
+and to unmap it::
+
+	dma_unmap_single(dev, dma_handle, size, direction);
+
+You should call dma_mapping_error() as dma_map_single() could fail and return
+error.  Doing so will ensure that the mapping code will work correctly on all
+DMA implementations without any dependency on the specifics of the underlying
+implementation. Using the returned address without checking for errors could
+result in failures ranging from panics to silent data corruption.  The same
+applies to dma_map_page() as well.
+
+You should call dma_unmap_single() when the DMA activity is finished, e.g.,
+from the interrupt which told you that the DMA transfer is done.
+
+Using CPU pointers like this for single mappings has a disadvantage:
+you cannot reference HIGHMEM memory in this way.  Thus, there is a
+map/unmap interface pair akin to dma_{map,unmap}_single().  These
+interfaces deal with page/offset pairs instead of CPU pointers.
+Specifically::
+
+	struct device *dev = &my_dev->dev;
+	dma_addr_t dma_handle;
+	struct page *page = buffer->page;
+	unsigned long offset = buffer->offset;
+	size_t size = buffer->len;
+
+	dma_handle = dma_map_page(dev, page, offset, size, direction);
+	if (dma_mapping_error(dev, dma_handle)) {
+		/*
+		 * reduce current DMA mapping usage,
+		 * delay and try again later or
+		 * reset driver.
+		 */
+		goto map_error_handling;
+	}
+
+	...
+
+	dma_unmap_page(dev, dma_handle, size, direction);
+
+Here, "offset" means byte offset within the given page.
+
+You should call dma_mapping_error() as dma_map_page() could fail and return
+error as outlined under the dma_map_single() discussion.
+
+You should call dma_unmap_page() when the DMA activity is finished, e.g.,
+from the interrupt which told you that the DMA transfer is done.
+
+With scatterlists, you map a region gathered from several regions by::
+
+	int i, count = dma_map_sg(dev, sglist, nents, direction);
+	struct scatterlist *sg;
+
+	for_each_sg(sglist, sg, count, i) {
+		hw_address[i] = sg_dma_address(sg);
+		hw_len[i] = sg_dma_len(sg);
+	}
+
+where nents is the number of entries in the sglist.
+
+The implementation is free to merge several consecutive sglist entries
+into one (e.g. if DMA mapping is done with PAGE_SIZE granularity, any
+consecutive sglist entries can be merged into one provided the first one
+ends and the second one starts on a page boundary - in fact this is a huge
+advantage for cards which either cannot do scatter-gather or have very
+limited number of scatter-gather entries) and returns the actual number
+of sg entries it mapped them to. On failure 0 is returned.
+
+Then you should loop count times (note: this can be less than nents times)
+and use sg_dma_address() and sg_dma_len() macros where you previously
+accessed sg->address and sg->length as shown above.
+
+To unmap a scatterlist, just call::
+
+	dma_unmap_sg(dev, sglist, nents, direction);
+
+Again, make sure DMA activity has already finished.
+
+.. note::
+
+	The 'nents' argument to the dma_unmap_sg call must be
+	the _same_ one you passed into the dma_map_sg call,
+	it should _NOT_ be the 'count' value _returned_ from the
+	dma_map_sg call.
+
+Every dma_map_{single,sg}() call should have its dma_unmap_{single,sg}()
+counterpart, because the DMA address space is a shared resource and
+you could render the machine unusable by consuming all DMA addresses.
+
+If you need to use the same streaming DMA region multiple times and touch
+the data in between the DMA transfers, the buffer needs to be synced
+properly in order for the CPU and device to see the most up-to-date and
+correct copy of the DMA buffer.
+
+So, firstly, just map it with dma_map_{single,sg}(), and after each DMA
+transfer call either::
+
+	dma_sync_single_for_cpu(dev, dma_handle, size, direction);
+
+or::
+
+	dma_sync_sg_for_cpu(dev, sglist, nents, direction);
+
+as appropriate.
+
+Then, if you wish to let the device get at the DMA area again,
+finish accessing the data with the CPU, and then before actually
+giving the buffer to the hardware call either::
+
+	dma_sync_single_for_device(dev, dma_handle, size, direction);
+
+or::
+
+	dma_sync_sg_for_device(dev, sglist, nents, direction);
+
+as appropriate.
+
+.. note::
+
+	      The 'nents' argument to dma_sync_sg_for_cpu() and
+	      dma_sync_sg_for_device() must be the same passed to
+	      dma_map_sg(). It is _NOT_ the count returned by
+	      dma_map_sg().
+
+After the last DMA transfer call one of the DMA unmap routines
+dma_unmap_{single,sg}(). If you don't touch the data from the first
+dma_map_*() call till dma_unmap_*(), then you don't have to call the
+dma_sync_*() routines at all.
+
+Here is pseudo code which shows a situation in which you would need
+to use the dma_sync_*() interfaces::
+
+	my_card_setup_receive_buffer(struct my_card *cp, char *buffer, int len)
+	{
+		dma_addr_t mapping;
+
+		mapping = dma_map_single(cp->dev, buffer, len, DMA_FROM_DEVICE);
+		if (dma_mapping_error(cp->dev, mapping)) {
+			/*
+			 * reduce current DMA mapping usage,
+			 * delay and try again later or
+			 * reset driver.
+			 */
+			goto map_error_handling;
+		}
+
+		cp->rx_buf = buffer;
+		cp->rx_len = len;
+		cp->rx_dma = mapping;
+
+		give_rx_buf_to_card(cp);
+	}
+
+	...
+
+	my_card_interrupt_handler(int irq, void *devid, struct pt_regs *regs)
+	{
+		struct my_card *cp = devid;
+
+		...
+		if (read_card_status(cp) == RX_BUF_TRANSFERRED) {
+			struct my_card_header *hp;
+
+			/* Examine the header to see if we wish
+			 * to accept the data.  But synchronize
+			 * the DMA transfer with the CPU first
+			 * so that we see updated contents.
+			 */
+			dma_sync_single_for_cpu(&cp->dev, cp->rx_dma,
+						cp->rx_len,
+						DMA_FROM_DEVICE);
+
+			/* Now it is safe to examine the buffer. */
+			hp = (struct my_card_header *) cp->rx_buf;
+			if (header_is_ok(hp)) {
+				dma_unmap_single(&cp->dev, cp->rx_dma, cp->rx_len,
+						 DMA_FROM_DEVICE);
+				pass_to_upper_layers(cp->rx_buf);
+				make_and_setup_new_rx_buf(cp);
+			} else {
+				/* CPU should not write to
+				 * DMA_FROM_DEVICE-mapped area,
+				 * so dma_sync_single_for_device() is
+				 * not needed here. It would be required
+				 * for DMA_BIDIRECTIONAL mapping if
+				 * the memory was modified.
+				 */
+				give_rx_buf_to_card(cp);
+			}
+		}
+	}
+
+Drivers converted fully to this interface should not use virt_to_bus() any
+longer, nor should they use bus_to_virt(). Some drivers have to be changed a
+little bit, because there is no longer an equivalent to bus_to_virt() in the
+dynamic DMA mapping scheme - you have to always store the DMA addresses
+returned by the dma_alloc_coherent(), dma_pool_alloc(), and dma_map_single()
+calls (dma_map_sg() stores them in the scatterlist itself if the platform
+supports dynamic DMA mapping in hardware) in your driver structures and/or
+in the card registers.
+
+All drivers should be using these interfaces with no exceptions.  It
+is planned to completely remove virt_to_bus() and bus_to_virt() as
+they are entirely deprecated.  Some ports already do not provide these
+as it is impossible to correctly support them.
+
+Handling Errors
+===============
+
+DMA address space is limited on some architectures and an allocation
+failure can be determined by:
+
+- checking if dma_alloc_coherent() returns NULL or dma_map_sg returns 0
+
+- checking the dma_addr_t returned from dma_map_single() and dma_map_page()
+  by using dma_mapping_error()::
+
+	dma_addr_t dma_handle;
+
+	dma_handle = dma_map_single(dev, addr, size, direction);
+	if (dma_mapping_error(dev, dma_handle)) {
+		/*
+		 * reduce current DMA mapping usage,
+		 * delay and try again later or
+		 * reset driver.
+		 */
+		goto map_error_handling;
+	}
+
+- unmap pages that are already mapped, when mapping error occurs in the middle
+  of a multiple page mapping attempt. These example are applicable to
+  dma_map_page() as well.
+
+Example 1::
+
+	dma_addr_t dma_handle1;
+	dma_addr_t dma_handle2;
+
+	dma_handle1 = dma_map_single(dev, addr, size, direction);
+	if (dma_mapping_error(dev, dma_handle1)) {
+		/*
+		 * reduce current DMA mapping usage,
+		 * delay and try again later or
+		 * reset driver.
+		 */
+		goto map_error_handling1;
+	}
+	dma_handle2 = dma_map_single(dev, addr, size, direction);
+	if (dma_mapping_error(dev, dma_handle2)) {
+		/*
+		 * reduce current DMA mapping usage,
+		 * delay and try again later or
+		 * reset driver.
+		 */
+		goto map_error_handling2;
+	}
+
+	...
+
+	map_error_handling2:
+		dma_unmap_single(dma_handle1);
+	map_error_handling1:
+
+Example 2::
+
+	/*
+	 * if buffers are allocated in a loop, unmap all mapped buffers when
+	 * mapping error is detected in the middle
+	 */
+
+	dma_addr_t dma_addr;
+	dma_addr_t array[DMA_BUFFERS];
+	int save_index = 0;
+
+	for (i = 0; i < DMA_BUFFERS; i++) {
+
+		...
+
+		dma_addr = dma_map_single(dev, addr, size, direction);
+		if (dma_mapping_error(dev, dma_addr)) {
+			/*
+			 * reduce current DMA mapping usage,
+			 * delay and try again later or
+			 * reset driver.
+			 */
+			goto map_error_handling;
+		}
+		array[i].dma_addr = dma_addr;
+		save_index++;
+	}
+
+	...
+
+	map_error_handling:
+
+	for (i = 0; i < save_index; i++) {
+
+		...
+
+		dma_unmap_single(array[i].dma_addr);
+	}
+
+Networking drivers must call dev_kfree_skb() to free the socket buffer
+and return NETDEV_TX_OK if the DMA mapping fails on the transmit hook
+(ndo_start_xmit). This means that the socket buffer is just dropped in
+the failure case.
+
+SCSI drivers must return SCSI_MLQUEUE_HOST_BUSY if the DMA mapping
+fails in the queuecommand hook. This means that the SCSI subsystem
+passes the command to the driver again later.
+
+Optimizing Unmap State Space Consumption
+========================================
+
+On many platforms, dma_unmap_{single,page}() is simply a nop.
+Therefore, keeping track of the mapping address and length is a waste
+of space.  Instead of filling your drivers up with ifdefs and the like
+to "work around" this (which would defeat the whole purpose of a
+portable API) the following facilities are provided.
+
+Actually, instead of describing the macros one by one, we'll
+transform some example code.
+
+1) Use DEFINE_DMA_UNMAP_{ADDR,LEN} in state saving structures.
+   Example, before::
+
+	struct ring_state {
+		struct sk_buff *skb;
+		dma_addr_t mapping;
+		__u32 len;
+	};
+
+   after::
+
+	struct ring_state {
+		struct sk_buff *skb;
+		DEFINE_DMA_UNMAP_ADDR(mapping);
+		DEFINE_DMA_UNMAP_LEN(len);
+	};
+
+2) Use dma_unmap_{addr,len}_set() to set these values.
+   Example, before::
+
+	ringp->mapping = FOO;
+	ringp->len = BAR;
+
+   after::
+
+	dma_unmap_addr_set(ringp, mapping, FOO);
+	dma_unmap_len_set(ringp, len, BAR);
+
+3) Use dma_unmap_{addr,len}() to access these values.
+   Example, before::
+
+	dma_unmap_single(dev, ringp->mapping, ringp->len,
+			 DMA_FROM_DEVICE);
+
+   after::
+
+	dma_unmap_single(dev,
+			 dma_unmap_addr(ringp, mapping),
+			 dma_unmap_len(ringp, len),
+			 DMA_FROM_DEVICE);
+
+It really should be self-explanatory.  We treat the ADDR and LEN
+separately, because it is possible for an implementation to only
+need the address in order to perform the unmap operation.
+
+Platform Issues
+===============
+
+If you are just writing drivers for Linux and do not maintain
+an architecture port for the kernel, you can safely skip down
+to "Closing".
+
+1) Struct scatterlist requirements.
+
+   You need to enable CONFIG_NEED_SG_DMA_LENGTH if the architecture
+   supports IOMMUs (including software IOMMU).
+
+2) ARCH_DMA_MINALIGN
+
+   Architectures must ensure that kmalloc'ed buffer is
+   DMA-safe. Drivers and subsystems depend on it. If an architecture
+   isn't fully DMA-coherent (i.e. hardware doesn't ensure that data in
+   the CPU cache is identical to data in main memory),
+   ARCH_DMA_MINALIGN must be set so that the memory allocator
+   makes sure that kmalloc'ed buffer doesn't share a cache line with
+   the others. See arch/arm/include/asm/cache.h as an example.
+
+   Note that ARCH_DMA_MINALIGN is about DMA memory alignment
+   constraints. You don't need to worry about the architecture data
+   alignment constraints (e.g. the alignment constraints about 64-bit
+   objects).
+
+Closing
+=======
+
+This document, and the API itself, would not be in its current
+form without the feedback and suggestions from numerous individuals.
+We would like to specifically mention, in no particular order, the
+following people::
+
+	Russell King <rmk@arm.linux.org.uk>
+	Leo Dagum <dagum@barrel.engr.sgi.com>
+	Ralf Baechle <ralf@oss.sgi.com>
+	Grant Grundler <grundler@cup.hp.com>
+	Jay Estabrook <Jay.Estabrook@compaq.com>
+	Thomas Sailer <sailer@ife.ee.ethz.ch>
+	Andrea Arcangeli <andrea@suse.de>
+	Jens Axboe <jens.axboe@oracle.com>
+	David Mosberger-Tang <davidm@hpl.hp.com>
diff --git a/Documentation/core-api/dma-api.rst b/Documentation/core-api/dma-api.rst
new file mode 100644
index 000000000..75cb757bb
--- /dev/null
+++ b/Documentation/core-api/dma-api.rst
@@ -0,0 +1,763 @@
+============================================
+Dynamic DMA mapping using the generic device
+============================================
+
+:Author: James E.J. Bottomley <James.Bottomley@HansenPartnership.com>
+
+This document describes the DMA API.  For a more gentle introduction
+of the API (and actual examples), see :doc:`/core-api/dma-api-howto`.
+
+This API is split into two pieces.  Part I describes the basic API.
+Part II describes extensions for supporting non-consistent memory
+machines.  Unless you know that your driver absolutely has to support
+non-consistent platforms (this is usually only legacy platforms) you
+should only use the API described in part I.
+
+Part I - dma_API
+----------------
+
+To get the dma_API, you must #include <linux/dma-mapping.h>.  This
+provides dma_addr_t and the interfaces described below.
+
+A dma_addr_t can hold any valid DMA address for the platform.  It can be
+given to a device to use as a DMA source or target.  A CPU cannot reference
+a dma_addr_t directly because there may be translation between its physical
+address space and the DMA address space.
+
+Part Ia - Using large DMA-coherent buffers
+------------------------------------------
+
+::
+
+	void *
+	dma_alloc_coherent(struct device *dev, size_t size,
+			   dma_addr_t *dma_handle, gfp_t flag)
+
+Consistent memory is memory for which a write by either the device or
+the processor can immediately be read by the processor or device
+without having to worry about caching effects.  (You may however need
+to make sure to flush the processor's write buffers before telling
+devices to read that memory.)
+
+This routine allocates a region of <size> bytes of consistent memory.
+
+It returns a pointer to the allocated region (in the processor's virtual
+address space) or NULL if the allocation failed.
+
+It also returns a <dma_handle> which may be cast to an unsigned integer the
+same width as the bus and given to the device as the DMA address base of
+the region.
+
+Note: consistent memory can be expensive on some platforms, and the
+minimum allocation length may be as big as a page, so you should
+consolidate your requests for consistent memory as much as possible.
+The simplest way to do that is to use the dma_pool calls (see below).
+
+The flag parameter (dma_alloc_coherent() only) allows the caller to
+specify the ``GFP_`` flags (see kmalloc()) for the allocation (the
+implementation may choose to ignore flags that affect the location of
+the returned memory, like GFP_DMA).
+
+::
+
+	void
+	dma_free_coherent(struct device *dev, size_t size, void *cpu_addr,
+			  dma_addr_t dma_handle)
+
+Free a region of consistent memory you previously allocated.  dev,
+size and dma_handle must all be the same as those passed into
+dma_alloc_coherent().  cpu_addr must be the virtual address returned by
+the dma_alloc_coherent().
+
+Note that unlike their sibling allocation calls, these routines
+may only be called with IRQs enabled.
+
+
+Part Ib - Using small DMA-coherent buffers
+------------------------------------------
+
+To get this part of the dma_API, you must #include <linux/dmapool.h>
+
+Many drivers need lots of small DMA-coherent memory regions for DMA
+descriptors or I/O buffers.  Rather than allocating in units of a page
+or more using dma_alloc_coherent(), you can use DMA pools.  These work
+much like a struct kmem_cache, except that they use the DMA-coherent allocator,
+not __get_free_pages().  Also, they understand common hardware constraints
+for alignment, like queue heads needing to be aligned on N-byte boundaries.
+
+
+::
+
+	struct dma_pool *
+	dma_pool_create(const char *name, struct device *dev,
+			size_t size, size_t align, size_t alloc);
+
+dma_pool_create() initializes a pool of DMA-coherent buffers
+for use with a given device.  It must be called in a context which
+can sleep.
+
+The "name" is for diagnostics (like a struct kmem_cache name); dev and size
+are like what you'd pass to dma_alloc_coherent().  The device's hardware
+alignment requirement for this type of data is "align" (which is expressed
+in bytes, and must be a power of two).  If your device has no boundary
+crossing restrictions, pass 0 for alloc; passing 4096 says memory allocated
+from this pool must not cross 4KByte boundaries.
+
+::
+
+	void *
+	dma_pool_zalloc(struct dma_pool *pool, gfp_t mem_flags,
+		        dma_addr_t *handle)
+
+Wraps dma_pool_alloc() and also zeroes the returned memory if the
+allocation attempt succeeded.
+
+
+::
+
+	void *
+	dma_pool_alloc(struct dma_pool *pool, gfp_t gfp_flags,
+		       dma_addr_t *dma_handle);
+
+This allocates memory from the pool; the returned memory will meet the
+size and alignment requirements specified at creation time.  Pass
+GFP_ATOMIC to prevent blocking, or if it's permitted (not
+in_interrupt, not holding SMP locks), pass GFP_KERNEL to allow
+blocking.  Like dma_alloc_coherent(), this returns two values:  an
+address usable by the CPU, and the DMA address usable by the pool's
+device.
+
+::
+
+	void
+	dma_pool_free(struct dma_pool *pool, void *vaddr,
+		      dma_addr_t addr);
+
+This puts memory back into the pool.  The pool is what was passed to
+dma_pool_alloc(); the CPU (vaddr) and DMA addresses are what
+were returned when that routine allocated the memory being freed.
+
+::
+
+	void
+	dma_pool_destroy(struct dma_pool *pool);
+
+dma_pool_destroy() frees the resources of the pool.  It must be
+called in a context which can sleep.  Make sure you've freed all allocated
+memory back to the pool before you destroy it.
+
+
+Part Ic - DMA addressing limitations
+------------------------------------
+
+::
+
+	int
+	dma_set_mask_and_coherent(struct device *dev, u64 mask)
+
+Checks to see if the mask is possible and updates the device
+streaming and coherent DMA mask parameters if it is.
+
+Returns: 0 if successful and a negative error if not.
+
+::
+
+	int
+	dma_set_mask(struct device *dev, u64 mask)
+
+Checks to see if the mask is possible and updates the device
+parameters if it is.
+
+Returns: 0 if successful and a negative error if not.
+
+::
+
+	int
+	dma_set_coherent_mask(struct device *dev, u64 mask)
+
+Checks to see if the mask is possible and updates the device
+parameters if it is.
+
+Returns: 0 if successful and a negative error if not.
+
+::
+
+	u64
+	dma_get_required_mask(struct device *dev)
+
+This API returns the mask that the platform requires to
+operate efficiently.  Usually this means the returned mask
+is the minimum required to cover all of memory.  Examining the
+required mask gives drivers with variable descriptor sizes the
+opportunity to use smaller descriptors as necessary.
+
+Requesting the required mask does not alter the current mask.  If you
+wish to take advantage of it, you should issue a dma_set_mask()
+call to set the mask to the value returned.
+
+::
+
+	size_t
+	dma_max_mapping_size(struct device *dev);
+
+Returns the maximum size of a mapping for the device. The size parameter
+of the mapping functions like dma_map_single(), dma_map_page() and
+others should not be larger than the returned value.
+
+::
+
+	bool
+	dma_need_sync(struct device *dev, dma_addr_t dma_addr);
+
+Returns %true if dma_sync_single_for_{device,cpu} calls are required to
+transfer memory ownership.  Returns %false if those calls can be skipped.
+
+::
+
+	unsigned long
+	dma_get_merge_boundary(struct device *dev);
+
+Returns the DMA merge boundary. If the device cannot merge any the DMA address
+segments, the function returns 0.
+
+Part Id - Streaming DMA mappings
+--------------------------------
+
+::
+
+	dma_addr_t
+	dma_map_single(struct device *dev, void *cpu_addr, size_t size,
+		       enum dma_data_direction direction)
+
+Maps a piece of processor virtual memory so it can be accessed by the
+device and returns the DMA address of the memory.
+
+The direction for both APIs may be converted freely by casting.
+However the dma_API uses a strongly typed enumerator for its
+direction:
+
+======================= =============================================
+DMA_NONE		no direction (used for debugging)
+DMA_TO_DEVICE		data is going from the memory to the device
+DMA_FROM_DEVICE		data is coming from the device to the memory
+DMA_BIDIRECTIONAL	direction isn't known
+======================= =============================================
+
+.. note::
+
+	Not all memory regions in a machine can be mapped by this API.
+	Further, contiguous kernel virtual space may not be contiguous as
+	physical memory.  Since this API does not provide any scatter/gather
+	capability, it will fail if the user tries to map a non-physically
+	contiguous piece of memory.  For this reason, memory to be mapped by
+	this API should be obtained from sources which guarantee it to be
+	physically contiguous (like kmalloc).
+
+	Further, the DMA address of the memory must be within the
+	dma_mask of the device (the dma_mask is a bit mask of the
+	addressable region for the device, i.e., if the DMA address of
+	the memory ANDed with the dma_mask is still equal to the DMA
+	address, then the device can perform DMA to the memory).  To
+	ensure that the memory allocated by kmalloc is within the dma_mask,
+	the driver may specify various platform-dependent flags to restrict
+	the DMA address range of the allocation (e.g., on x86, GFP_DMA
+	guarantees to be within the first 16MB of available DMA addresses,
+	as required by ISA devices).
+
+	Note also that the above constraints on physical contiguity and
+	dma_mask may not apply if the platform has an IOMMU (a device which
+	maps an I/O DMA address to a physical memory address).  However, to be
+	portable, device driver writers may *not* assume that such an IOMMU
+	exists.
+
+.. warning::
+
+	Memory coherency operates at a granularity called the cache
+	line width.  In order for memory mapped by this API to operate
+	correctly, the mapped region must begin exactly on a cache line
+	boundary and end exactly on one (to prevent two separately mapped
+	regions from sharing a single cache line).  Since the cache line size
+	may not be known at compile time, the API will not enforce this
+	requirement.  Therefore, it is recommended that driver writers who
+	don't take special care to determine the cache line size at run time
+	only map virtual regions that begin and end on page boundaries (which
+	are guaranteed also to be cache line boundaries).
+
+	DMA_TO_DEVICE synchronisation must be done after the last modification
+	of the memory region by the software and before it is handed off to
+	the device.  Once this primitive is used, memory covered by this
+	primitive should be treated as read-only by the device.  If the device
+	may write to it at any point, it should be DMA_BIDIRECTIONAL (see
+	below).
+
+	DMA_FROM_DEVICE synchronisation must be done before the driver
+	accesses data that may be changed by the device.  This memory should
+	be treated as read-only by the driver.  If the driver needs to write
+	to it at any point, it should be DMA_BIDIRECTIONAL (see below).
+
+	DMA_BIDIRECTIONAL requires special handling: it means that the driver
+	isn't sure if the memory was modified before being handed off to the
+	device and also isn't sure if the device will also modify it.  Thus,
+	you must always sync bidirectional memory twice: once before the
+	memory is handed off to the device (to make sure all memory changes
+	are flushed from the processor) and once before the data may be
+	accessed after being used by the device (to make sure any processor
+	cache lines are updated with data that the device may have changed).
+
+::
+
+	void
+	dma_unmap_single(struct device *dev, dma_addr_t dma_addr, size_t size,
+			 enum dma_data_direction direction)
+
+Unmaps the region previously mapped.  All the parameters passed in
+must be identical to those passed in (and returned) by the mapping
+API.
+
+::
+
+	dma_addr_t
+	dma_map_page(struct device *dev, struct page *page,
+		     unsigned long offset, size_t size,
+		     enum dma_data_direction direction)
+
+	void
+	dma_unmap_page(struct device *dev, dma_addr_t dma_address, size_t size,
+		       enum dma_data_direction direction)
+
+API for mapping and unmapping for pages.  All the notes and warnings
+for the other mapping APIs apply here.  Also, although the <offset>
+and <size> parameters are provided to do partial page mapping, it is
+recommended that you never use these unless you really know what the
+cache width is.
+
+::
+
+	dma_addr_t
+	dma_map_resource(struct device *dev, phys_addr_t phys_addr, size_t size,
+			 enum dma_data_direction dir, unsigned long attrs)
+
+	void
+	dma_unmap_resource(struct device *dev, dma_addr_t addr, size_t size,
+			   enum dma_data_direction dir, unsigned long attrs)
+
+API for mapping and unmapping for MMIO resources. All the notes and
+warnings for the other mapping APIs apply here. The API should only be
+used to map device MMIO resources, mapping of RAM is not permitted.
+
+::
+
+	int
+	dma_mapping_error(struct device *dev, dma_addr_t dma_addr)
+
+In some circumstances dma_map_single(), dma_map_page() and dma_map_resource()
+will fail to create a mapping. A driver can check for these errors by testing
+the returned DMA address with dma_mapping_error(). A non-zero return value
+means the mapping could not be created and the driver should take appropriate
+action (e.g. reduce current DMA mapping usage or delay and try again later).
+
+::
+
+	int
+	dma_map_sg(struct device *dev, struct scatterlist *sg,
+		   int nents, enum dma_data_direction direction)
+
+Returns: the number of DMA address segments mapped (this may be shorter
+than <nents> passed in if some elements of the scatter/gather list are
+physically or virtually adjacent and an IOMMU maps them with a single
+entry).
+
+Please note that the sg cannot be mapped again if it has been mapped once.
+The mapping process is allowed to destroy information in the sg.
+
+As with the other mapping interfaces, dma_map_sg() can fail. When it
+does, 0 is returned and a driver must take appropriate action. It is
+critical that the driver do something, in the case of a block driver
+aborting the request or even oopsing is better than doing nothing and
+corrupting the filesystem.
+
+With scatterlists, you use the resulting mapping like this::
+
+	int i, count = dma_map_sg(dev, sglist, nents, direction);
+	struct scatterlist *sg;
+
+	for_each_sg(sglist, sg, count, i) {
+		hw_address[i] = sg_dma_address(sg);
+		hw_len[i] = sg_dma_len(sg);
+	}
+
+where nents is the number of entries in the sglist.
+
+The implementation is free to merge several consecutive sglist entries
+into one (e.g. with an IOMMU, or if several pages just happen to be
+physically contiguous) and returns the actual number of sg entries it
+mapped them to. On failure 0, is returned.
+
+Then you should loop count times (note: this can be less than nents times)
+and use sg_dma_address() and sg_dma_len() macros where you previously
+accessed sg->address and sg->length as shown above.
+
+::
+
+	void
+	dma_unmap_sg(struct device *dev, struct scatterlist *sg,
+		     int nents, enum dma_data_direction direction)
+
+Unmap the previously mapped scatter/gather list.  All the parameters
+must be the same as those and passed in to the scatter/gather mapping
+API.
+
+Note: <nents> must be the number you passed in, *not* the number of
+DMA address entries returned.
+
+::
+
+	void
+	dma_sync_single_for_cpu(struct device *dev, dma_addr_t dma_handle,
+				size_t size,
+				enum dma_data_direction direction)
+
+	void
+	dma_sync_single_for_device(struct device *dev, dma_addr_t dma_handle,
+				   size_t size,
+				   enum dma_data_direction direction)
+
+	void
+	dma_sync_sg_for_cpu(struct device *dev, struct scatterlist *sg,
+			    int nents,
+			    enum dma_data_direction direction)
+
+	void
+	dma_sync_sg_for_device(struct device *dev, struct scatterlist *sg,
+			       int nents,
+			       enum dma_data_direction direction)
+
+Synchronise a single contiguous or scatter/gather mapping for the CPU
+and device. With the sync_sg API, all the parameters must be the same
+as those passed into the single mapping API. With the sync_single API,
+you can use dma_handle and size parameters that aren't identical to
+those passed into the single mapping API to do a partial sync.
+
+
+.. note::
+
+   You must do this:
+
+   - Before reading values that have been written by DMA from the device
+     (use the DMA_FROM_DEVICE direction)
+   - After writing values that will be written to the device using DMA
+     (use the DMA_TO_DEVICE) direction
+   - before *and* after handing memory to the device if the memory is
+     DMA_BIDIRECTIONAL
+
+See also dma_map_single().
+
+::
+
+	dma_addr_t
+	dma_map_single_attrs(struct device *dev, void *cpu_addr, size_t size,
+			     enum dma_data_direction dir,
+			     unsigned long attrs)
+
+	void
+	dma_unmap_single_attrs(struct device *dev, dma_addr_t dma_addr,
+			       size_t size, enum dma_data_direction dir,
+			       unsigned long attrs)
+
+	int
+	dma_map_sg_attrs(struct device *dev, struct scatterlist *sgl,
+			 int nents, enum dma_data_direction dir,
+			 unsigned long attrs)
+
+	void
+	dma_unmap_sg_attrs(struct device *dev, struct scatterlist *sgl,
+			   int nents, enum dma_data_direction dir,
+			   unsigned long attrs)
+
+The four functions above are just like the counterpart functions
+without the _attrs suffixes, except that they pass an optional
+dma_attrs.
+
+The interpretation of DMA attributes is architecture-specific, and
+each attribute should be documented in :doc:`/core-api/dma-attributes`.
+
+If dma_attrs are 0, the semantics of each of these functions
+is identical to those of the corresponding function
+without the _attrs suffix. As a result dma_map_single_attrs()
+can generally replace dma_map_single(), etc.
+
+As an example of the use of the ``*_attrs`` functions, here's how
+you could pass an attribute DMA_ATTR_FOO when mapping memory
+for DMA::
+
+	#include <linux/dma-mapping.h>
+	/* DMA_ATTR_FOO should be defined in linux/dma-mapping.h and
+	* documented in Documentation/core-api/dma-attributes.rst */
+	...
+
+		unsigned long attr;
+		attr |= DMA_ATTR_FOO;
+		....
+		n = dma_map_sg_attrs(dev, sg, nents, DMA_TO_DEVICE, attr);
+		....
+
+Architectures that care about DMA_ATTR_FOO would check for its
+presence in their implementations of the mapping and unmapping
+routines, e.g.:::
+
+	void whizco_dma_map_sg_attrs(struct device *dev, dma_addr_t dma_addr,
+				     size_t size, enum dma_data_direction dir,
+				     unsigned long attrs)
+	{
+		....
+		if (attrs & DMA_ATTR_FOO)
+			/* twizzle the frobnozzle */
+		....
+	}
+
+
+Part II - Non-coherent DMA allocations
+--------------------------------------
+
+These APIs allow to allocate pages that are guaranteed to be DMA addressable
+by the passed in device, but which need explicit management of memory ownership
+for the kernel vs the device.
+
+If you don't understand how cache line coherency works between a processor and
+an I/O device, you should not be using this part of the API.
+
+::
+
+	void *
+	dma_alloc_noncoherent(struct device *dev, size_t size,
+			dma_addr_t *dma_handle, enum dma_data_direction dir,
+			gfp_t gfp)
+
+This routine allocates a region of <size> bytes of consistent memory.  It
+returns a pointer to the allocated region (in the processor's virtual address
+space) or NULL if the allocation failed.  The returned memory may or may not
+be in the kernel direct mapping.  Drivers must not call virt_to_page on
+the returned memory region.
+
+It also returns a <dma_handle> which may be cast to an unsigned integer the
+same width as the bus and given to the device as the DMA address base of
+the region.
+
+The dir parameter specified if data is read and/or written by the device,
+see dma_map_single() for details.
+
+The gfp parameter allows the caller to specify the ``GFP_`` flags (see
+kmalloc()) for the allocation, but rejects flags used to specify a memory
+zone such as GFP_DMA or GFP_HIGHMEM.
+
+Before giving the memory to the device, dma_sync_single_for_device() needs
+to be called, and before reading memory written by the device,
+dma_sync_single_for_cpu(), just like for streaming DMA mappings that are
+reused.
+
+::
+
+	void
+	dma_free_noncoherent(struct device *dev, size_t size, void *cpu_addr,
+			dma_addr_t dma_handle, enum dma_data_direction dir)
+
+Free a region of memory previously allocated using dma_alloc_noncoherent().
+dev, size and dma_handle and dir must all be the same as those passed into
+dma_alloc_noncoherent().  cpu_addr must be the virtual address returned by
+dma_alloc_noncoherent().
+
+::
+
+	struct page *
+	dma_alloc_pages(struct device *dev, size_t size, dma_addr_t *dma_handle,
+			enum dma_data_direction dir, gfp_t gfp)
+
+This routine allocates a region of <size> bytes of non-coherent memory.  It
+returns a pointer to first struct page for the region, or NULL if the
+allocation failed. The resulting struct page can be used for everything a
+struct page is suitable for.
+
+It also returns a <dma_handle> which may be cast to an unsigned integer the
+same width as the bus and given to the device as the DMA address base of
+the region.
+
+The dir parameter specified if data is read and/or written by the device,
+see dma_map_single() for details.
+
+The gfp parameter allows the caller to specify the ``GFP_`` flags (see
+kmalloc()) for the allocation, but rejects flags used to specify a memory
+zone such as GFP_DMA or GFP_HIGHMEM.
+
+Before giving the memory to the device, dma_sync_single_for_device() needs
+to be called, and before reading memory written by the device,
+dma_sync_single_for_cpu(), just like for streaming DMA mappings that are
+reused.
+
+::
+
+	void
+	dma_free_pages(struct device *dev, size_t size, struct page *page,
+			dma_addr_t dma_handle, enum dma_data_direction dir)
+
+Free a region of memory previously allocated using dma_alloc_pages().
+dev, size and dma_handle and dir must all be the same as those passed into
+dma_alloc_noncoherent().  page must be the pointer returned by
+dma_alloc_pages().
+
+::
+
+	int
+	dma_get_cache_alignment(void)
+
+Returns the processor cache alignment.  This is the absolute minimum
+alignment *and* width that you must observe when either mapping
+memory or doing partial flushes.
+
+.. note::
+
+	This API may return a number *larger* than the actual cache
+	line, but it will guarantee that one or more cache lines fit exactly
+	into the width returned by this call.  It will also always be a power
+	of two for easy alignment.
+
+
+Part III - Debug drivers use of the DMA-API
+-------------------------------------------
+
+The DMA-API as described above has some constraints. DMA addresses must be
+released with the corresponding function with the same size for example. With
+the advent of hardware IOMMUs it becomes more and more important that drivers
+do not violate those constraints. In the worst case such a violation can
+result in data corruption up to destroyed filesystems.
+
+To debug drivers and find bugs in the usage of the DMA-API checking code can
+be compiled into the kernel which will tell the developer about those
+violations. If your architecture supports it you can select the "Enable
+debugging of DMA-API usage" option in your kernel configuration. Enabling this
+option has a performance impact. Do not enable it in production kernels.
+
+If you boot the resulting kernel will contain code which does some bookkeeping
+about what DMA memory was allocated for which device. If this code detects an
+error it prints a warning message with some details into your kernel log. An
+example warning message may look like this::
+
+	WARNING: at /data2/repos/linux-2.6-iommu/lib/dma-debug.c:448
+		check_unmap+0x203/0x490()
+	Hardware name:
+	forcedeth 0000:00:08.0: DMA-API: device driver frees DMA memory with wrong
+		function [device address=0x00000000640444be] [size=66 bytes] [mapped as
+	single] [unmapped as page]
+	Modules linked in: nfsd exportfs bridge stp llc r8169
+	Pid: 0, comm: swapper Tainted: G        W  2.6.28-dmatest-09289-g8bb99c0 #1
+	Call Trace:
+	<IRQ>  [<ffffffff80240b22>] warn_slowpath+0xf2/0x130
+	[<ffffffff80647b70>] _spin_unlock+0x10/0x30
+	[<ffffffff80537e75>] usb_hcd_link_urb_to_ep+0x75/0xc0
+	[<ffffffff80647c22>] _spin_unlock_irqrestore+0x12/0x40
+	[<ffffffff8055347f>] ohci_urb_enqueue+0x19f/0x7c0
+	[<ffffffff80252f96>] queue_work+0x56/0x60
+	[<ffffffff80237e10>] enqueue_task_fair+0x20/0x50
+	[<ffffffff80539279>] usb_hcd_submit_urb+0x379/0xbc0
+	[<ffffffff803b78c3>] cpumask_next_and+0x23/0x40
+	[<ffffffff80235177>] find_busiest_group+0x207/0x8a0
+	[<ffffffff8064784f>] _spin_lock_irqsave+0x1f/0x50
+	[<ffffffff803c7ea3>] check_unmap+0x203/0x490
+	[<ffffffff803c8259>] debug_dma_unmap_page+0x49/0x50
+	[<ffffffff80485f26>] nv_tx_done_optimized+0xc6/0x2c0
+	[<ffffffff80486c13>] nv_nic_irq_optimized+0x73/0x2b0
+	[<ffffffff8026df84>] handle_IRQ_event+0x34/0x70
+	[<ffffffff8026ffe9>] handle_edge_irq+0xc9/0x150
+	[<ffffffff8020e3ab>] do_IRQ+0xcb/0x1c0
+	[<ffffffff8020c093>] ret_from_intr+0x0/0xa
+	<EOI> <4>---[ end trace f6435a98e2a38c0e ]---
+
+The driver developer can find the driver and the device including a stacktrace
+of the DMA-API call which caused this warning.
+
+Per default only the first error will result in a warning message. All other
+errors will only silently counted. This limitation exist to prevent the code
+from flooding your kernel log. To support debugging a device driver this can
+be disabled via debugfs. See the debugfs interface documentation below for
+details.
+
+The debugfs directory for the DMA-API debugging code is called dma-api/. In
+this directory the following files can currently be found:
+
+=============================== ===============================================
+dma-api/all_errors		This file contains a numeric value. If this
+				value is not equal to zero the debugging code
+				will print a warning for every error it finds
+				into the kernel log. Be careful with this
+				option, as it can easily flood your logs.
+
+dma-api/disabled		This read-only file contains the character 'Y'
+				if the debugging code is disabled. This can
+				happen when it runs out of memory or if it was
+				disabled at boot time
+
+dma-api/dump			This read-only file contains current DMA
+				mappings.
+
+dma-api/error_count		This file is read-only and shows the total
+				numbers of errors found.
+
+dma-api/num_errors		The number in this file shows how many
+				warnings will be printed to the kernel log
+				before it stops. This number is initialized to
+				one at system boot and be set by writing into
+				this file
+
+dma-api/min_free_entries	This read-only file can be read to get the
+				minimum number of free dma_debug_entries the
+				allocator has ever seen. If this value goes
+				down to zero the code will attempt to increase
+				nr_total_entries to compensate.
+
+dma-api/num_free_entries	The current number of free dma_debug_entries
+				in the allocator.
+
+dma-api/nr_total_entries	The total number of dma_debug_entries in the
+				allocator, both free and used.
+
+dma-api/driver_filter		You can write a name of a driver into this file
+				to limit the debug output to requests from that
+				particular driver. Write an empty string to
+				that file to disable the filter and see
+				all errors again.
+=============================== ===============================================
+
+If you have this code compiled into your kernel it will be enabled by default.
+If you want to boot without the bookkeeping anyway you can provide
+'dma_debug=off' as a boot parameter. This will disable DMA-API debugging.
+Notice that you can not enable it again at runtime. You have to reboot to do
+so.
+
+If you want to see debug messages only for a special device driver you can
+specify the dma_debug_driver=<drivername> parameter. This will enable the
+driver filter at boot time. The debug code will only print errors for that
+driver afterwards. This filter can be disabled or changed later using debugfs.
+
+When the code disables itself at runtime this is most likely because it ran
+out of dma_debug_entries and was unable to allocate more on-demand. 65536
+entries are preallocated at boot - if this is too low for you boot with
+'dma_debug_entries=<your_desired_number>' to overwrite the default. Note
+that the code allocates entries in batches, so the exact number of
+preallocated entries may be greater than the actual number requested. The
+code will print to the kernel log each time it has dynamically allocated
+as many entries as were initially preallocated. This is to indicate that a
+larger preallocation size may be appropriate, or if it happens continually
+that a driver may be leaking mappings.
+
+::
+
+	void
+	debug_dma_mapping_error(struct device *dev, dma_addr_t dma_addr);
+
+dma-debug interface debug_dma_mapping_error() to debug drivers that fail
+to check DMA mapping errors on addresses returned by dma_map_single() and
+dma_map_page() interfaces. This interface clears a flag set by
+debug_dma_map_page() to indicate that dma_mapping_error() has been called by
+the driver. When driver does unmap, debug_dma_unmap() checks the flag and if
+this flag is still set, prints warning message that includes call trace that
+leads up to the unmap. This interface can be called from dma_mapping_error()
+routines to enable DMA mapping error check debugging.
diff --git a/Documentation/core-api/dma-attributes.rst b/Documentation/core-api/dma-attributes.rst
new file mode 100644
index 000000000..1887d92e8
--- /dev/null
+++ b/Documentation/core-api/dma-attributes.rst
@@ -0,0 +1,132 @@
+==============
+DMA attributes
+==============
+
+This document describes the semantics of the DMA attributes that are
+defined in linux/dma-mapping.h.
+
+DMA_ATTR_WEAK_ORDERING
+----------------------
+
+DMA_ATTR_WEAK_ORDERING specifies that reads and writes to the mapping
+may be weakly ordered, that is that reads and writes may pass each other.
+
+Since it is optional for platforms to implement DMA_ATTR_WEAK_ORDERING,
+those that do not will simply ignore the attribute and exhibit default
+behavior.
+
+DMA_ATTR_WRITE_COMBINE
+----------------------
+
+DMA_ATTR_WRITE_COMBINE specifies that writes to the mapping may be
+buffered to improve performance.
+
+Since it is optional for platforms to implement DMA_ATTR_WRITE_COMBINE,
+those that do not will simply ignore the attribute and exhibit default
+behavior.
+
+DMA_ATTR_NO_KERNEL_MAPPING
+--------------------------
+
+DMA_ATTR_NO_KERNEL_MAPPING lets the platform to avoid creating a kernel
+virtual mapping for the allocated buffer. On some architectures creating
+such mapping is non-trivial task and consumes very limited resources
+(like kernel virtual address space or dma consistent address space).
+Buffers allocated with this attribute can be only passed to user space
+by calling dma_mmap_attrs(). By using this API, you are guaranteeing
+that you won't dereference the pointer returned by dma_alloc_attr(). You
+can treat it as a cookie that must be passed to dma_mmap_attrs() and
+dma_free_attrs(). Make sure that both of these also get this attribute
+set on each call.
+
+Since it is optional for platforms to implement
+DMA_ATTR_NO_KERNEL_MAPPING, those that do not will simply ignore the
+attribute and exhibit default behavior.
+
+DMA_ATTR_SKIP_CPU_SYNC
+----------------------
+
+By default dma_map_{single,page,sg} functions family transfer a given
+buffer from CPU domain to device domain. Some advanced use cases might
+require sharing a buffer between more than one device. This requires
+having a mapping created separately for each device and is usually
+performed by calling dma_map_{single,page,sg} function more than once
+for the given buffer with device pointer to each device taking part in
+the buffer sharing. The first call transfers a buffer from 'CPU' domain
+to 'device' domain, what synchronizes CPU caches for the given region
+(usually it means that the cache has been flushed or invalidated
+depending on the dma direction). However, next calls to
+dma_map_{single,page,sg}() for other devices will perform exactly the
+same synchronization operation on the CPU cache. CPU cache synchronization
+might be a time consuming operation, especially if the buffers are
+large, so it is highly recommended to avoid it if possible.
+DMA_ATTR_SKIP_CPU_SYNC allows platform code to skip synchronization of
+the CPU cache for the given buffer assuming that it has been already
+transferred to 'device' domain. This attribute can be also used for
+dma_unmap_{single,page,sg} functions family to force buffer to stay in
+device domain after releasing a mapping for it. Use this attribute with
+care!
+
+DMA_ATTR_FORCE_CONTIGUOUS
+-------------------------
+
+By default DMA-mapping subsystem is allowed to assemble the buffer
+allocated by dma_alloc_attrs() function from individual pages if it can
+be mapped as contiguous chunk into device dma address space. By
+specifying this attribute the allocated buffer is forced to be contiguous
+also in physical memory.
+
+DMA_ATTR_ALLOC_SINGLE_PAGES
+---------------------------
+
+This is a hint to the DMA-mapping subsystem that it's probably not worth
+the time to try to allocate memory to in a way that gives better TLB
+efficiency (AKA it's not worth trying to build the mapping out of larger
+pages).  You might want to specify this if:
+
+- You know that the accesses to this memory won't thrash the TLB.
+  You might know that the accesses are likely to be sequential or
+  that they aren't sequential but it's unlikely you'll ping-pong
+  between many addresses that are likely to be in different physical
+  pages.
+- You know that the penalty of TLB misses while accessing the
+  memory will be small enough to be inconsequential.  If you are
+  doing a heavy operation like decryption or decompression this
+  might be the case.
+- You know that the DMA mapping is fairly transitory.  If you expect
+  the mapping to have a short lifetime then it may be worth it to
+  optimize allocation (avoid coming up with large pages) instead of
+  getting the slight performance win of larger pages.
+
+Setting this hint doesn't guarantee that you won't get huge pages, but it
+means that we won't try quite as hard to get them.
+
+.. note:: At the moment DMA_ATTR_ALLOC_SINGLE_PAGES is only implemented on ARM,
+	  though ARM64 patches will likely be posted soon.
+
+DMA_ATTR_NO_WARN
+----------------
+
+This tells the DMA-mapping subsystem to suppress allocation failure reports
+(similarly to __GFP_NOWARN).
+
+On some architectures allocation failures are reported with error messages
+to the system logs.  Although this can help to identify and debug problems,
+drivers which handle failures (eg, retry later) have no problems with them,
+and can actually flood the system logs with error messages that aren't any
+problem at all, depending on the implementation of the retry mechanism.
+
+So, this provides a way for drivers to avoid those error messages on calls
+where allocation failures are not a problem, and shouldn't bother the logs.
+
+.. note:: At the moment DMA_ATTR_NO_WARN is only implemented on PowerPC.
+
+DMA_ATTR_PRIVILEGED
+-------------------
+
+Some advanced peripherals such as remote processors and GPUs perform
+accesses to DMA buffers in both privileged "supervisor" and unprivileged
+"user" modes.  This attribute is used to indicate to the DMA-mapping
+subsystem that the buffer is fully accessible at the elevated privilege
+level (and ideally inaccessible or at least read-only at the
+lesser-privileged levels).
diff --git a/Documentation/core-api/dma-isa-lpc.rst b/Documentation/core-api/dma-isa-lpc.rst
new file mode 100644
index 000000000..e59a3d35a
--- /dev/null
+++ b/Documentation/core-api/dma-isa-lpc.rst
@@ -0,0 +1,152 @@
+============================
+DMA with ISA and LPC devices
+============================
+
+:Author: Pierre Ossman <drzeus@drzeus.cx>
+
+This document describes how to do DMA transfers using the old ISA DMA
+controller. Even though ISA is more or less dead today the LPC bus
+uses the same DMA system so it will be around for quite some time.
+
+Headers and dependencies
+------------------------
+
+To do ISA style DMA you need to include two headers::
+
+	#include <linux/dma-mapping.h>
+	#include <asm/dma.h>
+
+The first is the generic DMA API used to convert virtual addresses to
+bus addresses (see :doc:`/core-api/dma-api` for details).
+
+The second contains the routines specific to ISA DMA transfers. Since
+this is not present on all platforms make sure you construct your
+Kconfig to be dependent on ISA_DMA_API (not ISA) so that nobody tries
+to build your driver on unsupported platforms.
+
+Buffer allocation
+-----------------
+
+The ISA DMA controller has some very strict requirements on which
+memory it can access so extra care must be taken when allocating
+buffers.
+
+(You usually need a special buffer for DMA transfers instead of
+transferring directly to and from your normal data structures.)
+
+The DMA-able address space is the lowest 16 MB of _physical_ memory.
+Also the transfer block may not cross page boundaries (which are 64
+or 128 KiB depending on which channel you use).
+
+In order to allocate a piece of memory that satisfies all these
+requirements you pass the flag GFP_DMA to kmalloc.
+
+Unfortunately the memory available for ISA DMA is scarce so unless you
+allocate the memory during boot-up it's a good idea to also pass
+__GFP_RETRY_MAYFAIL and __GFP_NOWARN to make the allocator try a bit harder.
+
+(This scarcity also means that you should allocate the buffer as
+early as possible and not release it until the driver is unloaded.)
+
+Address translation
+-------------------
+
+To translate the virtual address to a bus address, use the normal DMA
+API. Do _not_ use isa_virt_to_bus() even though it does the same
+thing. The reason for this is that the function isa_virt_to_bus()
+will require a Kconfig dependency to ISA, not just ISA_DMA_API which
+is really all you need. Remember that even though the DMA controller
+has its origins in ISA it is used elsewhere.
+
+Note: x86_64 had a broken DMA API when it came to ISA but has since
+been fixed. If your arch has problems then fix the DMA API instead of
+reverting to the ISA functions.
+
+Channels
+--------
+
+A normal ISA DMA controller has 8 channels. The lower four are for
+8-bit transfers and the upper four are for 16-bit transfers.
+
+(Actually the DMA controller is really two separate controllers where
+channel 4 is used to give DMA access for the second controller (0-3).
+This means that of the four 16-bits channels only three are usable.)
+
+You allocate these in a similar fashion as all basic resources:
+
+extern int request_dma(unsigned int dmanr, const char * device_id);
+extern void free_dma(unsigned int dmanr);
+
+The ability to use 16-bit or 8-bit transfers is _not_ up to you as a
+driver author but depends on what the hardware supports. Check your
+specs or test different channels.
+
+Transfer data
+-------------
+
+Now for the good stuff, the actual DMA transfer. :)
+
+Before you use any ISA DMA routines you need to claim the DMA lock
+using claim_dma_lock(). The reason is that some DMA operations are
+not atomic so only one driver may fiddle with the registers at a
+time.
+
+The first time you use the DMA controller you should call
+clear_dma_ff(). This clears an internal register in the DMA
+controller that is used for the non-atomic operations. As long as you
+(and everyone else) uses the locking functions then you only need to
+reset this once.
+
+Next, you tell the controller in which direction you intend to do the
+transfer using set_dma_mode(). Currently you have the options
+DMA_MODE_READ and DMA_MODE_WRITE.
+
+Set the address from where the transfer should start (this needs to
+be 16-bit aligned for 16-bit transfers) and how many bytes to
+transfer. Note that it's _bytes_. The DMA routines will do all the
+required translation to values that the DMA controller understands.
+
+The final step is enabling the DMA channel and releasing the DMA
+lock.
+
+Once the DMA transfer is finished (or timed out) you should disable
+the channel again. You should also check get_dma_residue() to make
+sure that all data has been transferred.
+
+Example::
+
+	int flags, residue;
+
+	flags = claim_dma_lock();
+
+	clear_dma_ff();
+
+	set_dma_mode(channel, DMA_MODE_WRITE);
+	set_dma_addr(channel, phys_addr);
+	set_dma_count(channel, num_bytes);
+
+	dma_enable(channel);
+
+	release_dma_lock(flags);
+
+	while (!device_done());
+
+	flags = claim_dma_lock();
+
+	dma_disable(channel);
+
+	residue = dma_get_residue(channel);
+	if (residue != 0)
+		printk(KERN_ERR "driver: Incomplete DMA transfer!"
+			" %d bytes left!\n", residue);
+
+	release_dma_lock(flags);
+
+Suspend/resume
+--------------
+
+It is the driver's responsibility to make sure that the machine isn't
+suspended while a DMA transfer is in progress. Also, all DMA settings
+are lost when the system suspends so if your driver relies on the DMA
+controller being in a certain state then you have to restore these
+registers upon resume.
diff --git a/Documentation/core-api/errseq.rst b/Documentation/core-api/errseq.rst
new file mode 100644
index 000000000..ff332e272
--- /dev/null
+++ b/Documentation/core-api/errseq.rst
@@ -0,0 +1,159 @@
+=====================
+The errseq_t datatype
+=====================
+
+An errseq_t is a way of recording errors in one place, and allowing any
+number of "subscribers" to tell whether it has changed since a previous
+point where it was sampled.
+
+The initial use case for this is tracking errors for file
+synchronization syscalls (fsync, fdatasync, msync and sync_file_range),
+but it may be usable in other situations.
+
+It's implemented as an unsigned 32-bit value.  The low order bits are
+designated to hold an error code (between 1 and MAX_ERRNO).  The upper bits
+are used as a counter.  This is done with atomics instead of locking so that
+these functions can be called from any context.
+
+Note that there is a risk of collisions if new errors are being recorded
+frequently, since we have so few bits to use as a counter.
+
+To mitigate this, the bit between the error value and counter is used as
+a flag to tell whether the value has been sampled since a new value was
+recorded.  That allows us to avoid bumping the counter if no one has
+sampled it since the last time an error was recorded.
+
+Thus we end up with a value that looks something like this:
+
++--------------------------------------+----+------------------------+
+| 31..13                               | 12 | 11..0                  |
++--------------------------------------+----+------------------------+
+| counter                              | SF | errno                  |
++--------------------------------------+----+------------------------+
+
+The general idea is for "watchers" to sample an errseq_t value and keep
+it as a running cursor.  That value can later be used to tell whether
+any new errors have occurred since that sampling was done, and atomically
+record the state at the time that it was checked.  This allows us to
+record errors in one place, and then have a number of "watchers" that
+can tell whether the value has changed since they last checked it.
+
+A new errseq_t should always be zeroed out.  An errseq_t value of all zeroes
+is the special (but common) case where there has never been an error. An all
+zero value thus serves as the "epoch" if one wishes to know whether there
+has ever been an error set since it was first initialized.
+
+API usage
+=========
+
+Let me tell you a story about a worker drone.  Now, he's a good worker
+overall, but the company is a little...management heavy.  He has to
+report to 77 supervisors today, and tomorrow the "big boss" is coming in
+from out of town and he's sure to test the poor fellow too.
+
+They're all handing him work to do -- so much he can't keep track of who
+handed him what, but that's not really a big problem.  The supervisors
+just want to know when he's finished all of the work they've handed him so
+far and whether he made any mistakes since they last asked.
+
+He might have made the mistake on work they didn't actually hand him,
+but he can't keep track of things at that level of detail, all he can
+remember is the most recent mistake that he made.
+
+Here's our worker_drone representation::
+
+        struct worker_drone {
+                errseq_t        wd_err; /* for recording errors */
+        };
+
+Every day, the worker_drone starts out with a blank slate::
+
+        struct worker_drone wd;
+
+        wd.wd_err = (errseq_t)0;
+
+The supervisors come in and get an initial read for the day.  They
+don't care about anything that happened before their watch begins::
+
+        struct supervisor {
+                errseq_t        s_wd_err; /* private "cursor" for wd_err */
+                spinlock_t      s_wd_err_lock; /* protects s_wd_err */
+        }
+
+        struct supervisor       su;
+
+        su.s_wd_err = errseq_sample(&wd.wd_err);
+        spin_lock_init(&su.s_wd_err_lock);
+
+Now they start handing him tasks to do.  Every few minutes they ask him to
+finish up all of the work they've handed him so far.  Then they ask him
+whether he made any mistakes on any of it::
+
+        spin_lock(&su.su_wd_err_lock);
+        err = errseq_check_and_advance(&wd.wd_err, &su.s_wd_err);
+        spin_unlock(&su.su_wd_err_lock);
+
+Up to this point, that just keeps returning 0.
+
+Now, the owners of this company are quite miserly and have given him
+substandard equipment with which to do his job. Occasionally it
+glitches and he makes a mistake.  He sighs a heavy sigh, and marks it
+down::
+
+        errseq_set(&wd.wd_err, -EIO);
+
+...and then gets back to work.  The supervisors eventually poll again
+and they each get the error when they next check.  Subsequent calls will
+return 0, until another error is recorded, at which point it's reported
+to each of them once.
+
+Note that the supervisors can't tell how many mistakes he made, only
+whether one was made since they last checked, and the latest value
+recorded.
+
+Occasionally the big boss comes in for a spot check and asks the worker
+to do a one-off job for him. He's not really watching the worker
+full-time like the supervisors, but he does need to know whether a
+mistake occurred while his job was processing.
+
+He can just sample the current errseq_t in the worker, and then use that
+to tell whether an error has occurred later::
+
+        errseq_t since = errseq_sample(&wd.wd_err);
+        /* submit some work and wait for it to complete */
+        err = errseq_check(&wd.wd_err, since);
+
+Since he's just going to discard "since" after that point, he doesn't
+need to advance it here. He also doesn't need any locking since it's
+not usable by anyone else.
+
+Serializing errseq_t cursor updates
+===================================
+
+Note that the errseq_t API does not protect the errseq_t cursor during a
+check_and_advance_operation. Only the canonical error code is handled
+atomically.  In a situation where more than one task might be using the
+same errseq_t cursor at the same time, it's important to serialize
+updates to that cursor.
+
+If that's not done, then it's possible for the cursor to go backward
+in which case the same error could be reported more than once.
+
+Because of this, it's often advantageous to first do an errseq_check to
+see if anything has changed, and only later do an
+errseq_check_and_advance after taking the lock. e.g.::
+
+        if (errseq_check(&wd.wd_err, READ_ONCE(su.s_wd_err)) {
+                /* su.s_wd_err is protected by s_wd_err_lock */
+                spin_lock(&su.s_wd_err_lock);
+                err = errseq_check_and_advance(&wd.wd_err, &su.s_wd_err);
+                spin_unlock(&su.s_wd_err_lock);
+        }
+
+That avoids the spinlock in the common case where nothing has changed
+since the last time it was checked.
+
+Functions
+=========
+
+.. kernel-doc:: lib/errseq.c
diff --git a/Documentation/core-api/genalloc.rst b/Documentation/core-api/genalloc.rst
new file mode 100644
index 000000000..a5af2cbf5
--- /dev/null
+++ b/Documentation/core-api/genalloc.rst
@@ -0,0 +1,144 @@
+The genalloc/genpool subsystem
+==============================
+
+There are a number of memory-allocation subsystems in the kernel, each
+aimed at a specific need.  Sometimes, however, a kernel developer needs to
+implement a new allocator for a specific range of special-purpose memory;
+often that memory is located on a device somewhere.  The author of the
+driver for that device can certainly write a little allocator to get the
+job done, but that is the way to fill the kernel with dozens of poorly
+tested allocators.  Back in 2005, Jes Sorensen lifted one of those
+allocators from the sym53c8xx_2 driver and posted_ it as a generic module
+for the creation of ad hoc memory allocators.  This code was merged
+for the 2.6.13 release; it has been modified considerably since then.
+
+.. _posted: https://lwn.net/Articles/125842/
+
+Code using this allocator should include <linux/genalloc.h>.  The action
+begins with the creation of a pool using one of:
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_create		
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: devm_gen_pool_create
+
+A call to gen_pool_create() will create a pool.  The granularity of
+allocations is set with min_alloc_order; it is a log-base-2 number like
+those used by the page allocator, but it refers to bytes rather than pages.
+So, if min_alloc_order is passed as 3, then all allocations will be a
+multiple of eight bytes.  Increasing min_alloc_order decreases the memory
+required to track the memory in the pool.  The nid parameter specifies
+which NUMA node should be used for the allocation of the housekeeping
+structures; it can be -1 if the caller doesn't care.
+
+The "managed" interface devm_gen_pool_create() ties the pool to a
+specific device.  Among other things, it will automatically clean up the
+pool when the given device is destroyed.
+
+A pool is shut down with:
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_destroy
+
+It's worth noting that, if there are still allocations outstanding from the
+given pool, this function will take the rather extreme step of invoking
+BUG(), crashing the entire system.  You have been warned.
+
+A freshly created pool has no memory to allocate.  It is fairly useless in
+that state, so one of the first orders of business is usually to add memory
+to the pool.  That can be done with one of:
+
+.. kernel-doc:: include/linux/genalloc.h
+   :functions: gen_pool_add
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_add_owner
+
+A call to gen_pool_add() will place the size bytes of memory
+starting at addr (in the kernel's virtual address space) into the given
+pool, once again using nid as the node ID for ancillary memory allocations.
+The gen_pool_add_virt() variant associates an explicit physical
+address with the memory; this is only necessary if the pool will be used
+for DMA allocations.
+
+The functions for allocating memory from the pool (and putting it back)
+are:
+
+.. kernel-doc:: include/linux/genalloc.h
+   :functions: gen_pool_alloc
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_dma_alloc
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_free_owner
+
+As one would expect, gen_pool_alloc() will allocate size< bytes
+from the given pool.  The gen_pool_dma_alloc() variant allocates
+memory for use with DMA operations, returning the associated physical
+address in the space pointed to by dma.  This will only work if the memory
+was added with gen_pool_add_virt().  Note that this function
+departs from the usual genpool pattern of using unsigned long values to
+represent kernel addresses; it returns a void * instead.
+
+That all seems relatively simple; indeed, some developers clearly found it
+to be too simple.  After all, the interface above provides no control over
+how the allocation functions choose which specific piece of memory to
+return.  If that sort of control is needed, the following functions will be
+of interest:
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_alloc_algo_owner
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_set_algo
+
+Allocations with gen_pool_alloc_algo() specify an algorithm to be
+used to choose the memory to be allocated; the default algorithm can be set
+with gen_pool_set_algo().  The data value is passed to the
+algorithm; most ignore it, but it is occasionally needed.  One can,
+naturally, write a special-purpose algorithm, but there is a fair set
+already available:
+
+- gen_pool_first_fit is a simple first-fit allocator; this is the default
+  algorithm if none other has been specified.
+
+- gen_pool_first_fit_align forces the allocation to have a specific
+  alignment (passed via data in a genpool_data_align structure).
+
+- gen_pool_first_fit_order_align aligns the allocation to the order of the
+  size.  A 60-byte allocation will thus be 64-byte aligned, for example.
+
+- gen_pool_best_fit, as one would expect, is a simple best-fit allocator.
+
+- gen_pool_fixed_alloc allocates at a specific offset (passed in a
+  genpool_data_fixed structure via the data parameter) within the pool.
+  If the indicated memory is not available the allocation fails.
+
+There is a handful of other functions, mostly for purposes like querying
+the space available in the pool or iterating through chunks of memory.
+Most users, however, should not need much beyond what has been described
+above.  With luck, wider awareness of this module will help to prevent the
+writing of special-purpose memory allocators in the future.
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_virt_to_phys
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_for_each_chunk
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_has_addr
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_avail
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_size
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: gen_pool_get
+
+.. kernel-doc:: lib/genalloc.c
+   :functions: of_gen_pool_get
diff --git a/Documentation/core-api/generic-radix-tree.rst b/Documentation/core-api/generic-radix-tree.rst
new file mode 100644
index 000000000..ed42839ae
--- /dev/null
+++ b/Documentation/core-api/generic-radix-tree.rst
@@ -0,0 +1,12 @@
+=================================
+Generic radix trees/sparse arrays
+=================================
+
+.. kernel-doc:: include/linux/generic-radix-tree.h
+   :doc: Generic radix trees/sparse arrays
+
+generic radix tree functions
+----------------------------
+
+.. kernel-doc:: include/linux/generic-radix-tree.h
+   :functions:
diff --git a/Documentation/core-api/genericirq.rst b/Documentation/core-api/genericirq.rst
new file mode 100644
index 000000000..f959c9b53
--- /dev/null
+++ b/Documentation/core-api/genericirq.rst
@@ -0,0 +1,444 @@
+.. include:: <isonum.txt>
+
+==========================
+Linux generic IRQ handling
+==========================
+
+:Copyright: |copy| 2005-2010: Thomas Gleixner
+:Copyright: |copy| 2005-2006:  Ingo Molnar
+
+Introduction
+============
+
+The generic interrupt handling layer is designed to provide a complete
+abstraction of interrupt handling for device drivers. It is able to
+handle all the different types of interrupt controller hardware. Device
+drivers use generic API functions to request, enable, disable and free
+interrupts. The drivers do not have to know anything about interrupt
+hardware details, so they can be used on different platforms without
+code changes.
+
+This documentation is provided to developers who want to implement an
+interrupt subsystem based for their architecture, with the help of the
+generic IRQ handling layer.
+
+Rationale
+=========
+
+The original implementation of interrupt handling in Linux uses the
+__do_IRQ() super-handler, which is able to deal with every type of
+interrupt logic.
+
+Originally, Russell King identified different types of handlers to build
+a quite universal set for the ARM interrupt handler implementation in
+Linux 2.5/2.6. He distinguished between:
+
+-  Level type
+
+-  Edge type
+
+-  Simple type
+
+During the implementation we identified another type:
+
+-  Fast EOI type
+
+In the SMP world of the __do_IRQ() super-handler another type was
+identified:
+
+-  Per CPU type
+
+This split implementation of high-level IRQ handlers allows us to
+optimize the flow of the interrupt handling for each specific interrupt
+type. This reduces complexity in that particular code path and allows
+the optimized handling of a given type.
+
+The original general IRQ implementation used hw_interrupt_type
+structures and their ``->ack``, ``->end`` [etc.] callbacks to differentiate
+the flow control in the super-handler. This leads to a mix of flow logic
+and low-level hardware logic, and it also leads to unnecessary code
+duplication: for example in i386, there is an ``ioapic_level_irq`` and an
+``ioapic_edge_irq`` IRQ-type which share many of the low-level details but
+have different flow handling.
+
+A more natural abstraction is the clean separation of the 'irq flow' and
+the 'chip details'.
+
+Analysing a couple of architecture's IRQ subsystem implementations
+reveals that most of them can use a generic set of 'irq flow' methods
+and only need to add the chip-level specific code. The separation is
+also valuable for (sub)architectures which need specific quirks in the
+IRQ flow itself but not in the chip details - and thus provides a more
+transparent IRQ subsystem design.
+
+Each interrupt descriptor is assigned its own high-level flow handler,
+which is normally one of the generic implementations. (This high-level
+flow handler implementation also makes it simple to provide
+demultiplexing handlers which can be found in embedded platforms on
+various architectures.)
+
+The separation makes the generic interrupt handling layer more flexible
+and extensible. For example, an (sub)architecture can use a generic
+IRQ-flow implementation for 'level type' interrupts and add a
+(sub)architecture specific 'edge type' implementation.
+
+To make the transition to the new model easier and prevent the breakage
+of existing implementations, the __do_IRQ() super-handler is still
+available. This leads to a kind of duality for the time being. Over time
+the new model should be used in more and more architectures, as it
+enables smaller and cleaner IRQ subsystems. It's deprecated for three
+years now and about to be removed.
+
+Known Bugs And Assumptions
+==========================
+
+None (knock on wood).
+
+Abstraction layers
+==================
+
+There are three main levels of abstraction in the interrupt code:
+
+1. High-level driver API
+
+2. High-level IRQ flow handlers
+
+3. Chip-level hardware encapsulation
+
+Interrupt control flow
+----------------------
+
+Each interrupt is described by an interrupt descriptor structure
+irq_desc. The interrupt is referenced by an 'unsigned int' numeric
+value which selects the corresponding interrupt description structure in
+the descriptor structures array. The descriptor structure contains
+status information and pointers to the interrupt flow method and the
+interrupt chip structure which are assigned to this interrupt.
+
+Whenever an interrupt triggers, the low-level architecture code calls
+into the generic interrupt code by calling desc->handle_irq(). This
+high-level IRQ handling function only uses desc->irq_data.chip
+primitives referenced by the assigned chip descriptor structure.
+
+High-level Driver API
+---------------------
+
+The high-level Driver API consists of following functions:
+
+-  request_irq()
+
+-  request_threaded_irq()
+
+-  free_irq()
+
+-  disable_irq()
+
+-  enable_irq()
+
+-  disable_irq_nosync() (SMP only)
+
+-  synchronize_irq() (SMP only)
+
+-  irq_set_irq_type()
+
+-  irq_set_irq_wake()
+
+-  irq_set_handler_data()
+
+-  irq_set_chip()
+
+-  irq_set_chip_data()
+
+See the autogenerated function documentation for details.
+
+High-level IRQ flow handlers
+----------------------------
+
+The generic layer provides a set of pre-defined irq-flow methods:
+
+-  handle_level_irq()
+
+-  handle_edge_irq()
+
+-  handle_fasteoi_irq()
+
+-  handle_simple_irq()
+
+-  handle_percpu_irq()
+
+-  handle_edge_eoi_irq()
+
+-  handle_bad_irq()
+
+The interrupt flow handlers (either pre-defined or architecture
+specific) are assigned to specific interrupts by the architecture either
+during bootup or during device initialization.
+
+Default flow implementations
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+Helper functions
+^^^^^^^^^^^^^^^^
+
+The helper functions call the chip primitives and are used by the
+default flow implementations. The following helper functions are
+implemented (simplified excerpt)::
+
+    default_enable(struct irq_data *data)
+    {
+        desc->irq_data.chip->irq_unmask(data);
+    }
+
+    default_disable(struct irq_data *data)
+    {
+        if (!delay_disable(data))
+            desc->irq_data.chip->irq_mask(data);
+    }
+
+    default_ack(struct irq_data *data)
+    {
+        chip->irq_ack(data);
+    }
+
+    default_mask_ack(struct irq_data *data)
+    {
+        if (chip->irq_mask_ack) {
+            chip->irq_mask_ack(data);
+        } else {
+            chip->irq_mask(data);
+            chip->irq_ack(data);
+        }
+    }
+
+    noop(struct irq_data *data))
+    {
+    }
+
+
+
+Default flow handler implementations
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+Default Level IRQ flow handler
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+handle_level_irq provides a generic implementation for level-triggered
+interrupts.
+
+The following control flow is implemented (simplified excerpt)::
+
+    desc->irq_data.chip->irq_mask_ack();
+    handle_irq_event(desc->action);
+    desc->irq_data.chip->irq_unmask();
+
+
+Default Fast EOI IRQ flow handler
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+handle_fasteoi_irq provides a generic implementation for interrupts,
+which only need an EOI at the end of the handler.
+
+The following control flow is implemented (simplified excerpt)::
+
+    handle_irq_event(desc->action);
+    desc->irq_data.chip->irq_eoi();
+
+
+Default Edge IRQ flow handler
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+handle_edge_irq provides a generic implementation for edge-triggered
+interrupts.
+
+The following control flow is implemented (simplified excerpt)::
+
+    if (desc->status & running) {
+        desc->irq_data.chip->irq_mask_ack();
+        desc->status |= pending | masked;
+        return;
+    }
+    desc->irq_data.chip->irq_ack();
+    desc->status |= running;
+    do {
+        if (desc->status & masked)
+            desc->irq_data.chip->irq_unmask();
+        desc->status &= ~pending;
+        handle_irq_event(desc->action);
+    } while (status & pending);
+    desc->status &= ~running;
+
+
+Default simple IRQ flow handler
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+handle_simple_irq provides a generic implementation for simple
+interrupts.
+
+.. note::
+
+   The simple flow handler does not call any handler/chip primitives.
+
+The following control flow is implemented (simplified excerpt)::
+
+    handle_irq_event(desc->action);
+
+
+Default per CPU flow handler
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+handle_percpu_irq provides a generic implementation for per CPU
+interrupts.
+
+Per CPU interrupts are only available on SMP and the handler provides a
+simplified version without locking.
+
+The following control flow is implemented (simplified excerpt)::
+
+    if (desc->irq_data.chip->irq_ack)
+        desc->irq_data.chip->irq_ack();
+    handle_irq_event(desc->action);
+    if (desc->irq_data.chip->irq_eoi)
+        desc->irq_data.chip->irq_eoi();
+
+
+EOI Edge IRQ flow handler
+^^^^^^^^^^^^^^^^^^^^^^^^^
+
+handle_edge_eoi_irq provides an abnomination of the edge handler
+which is solely used to tame a badly wreckaged irq controller on
+powerpc/cell.
+
+Bad IRQ flow handler
+^^^^^^^^^^^^^^^^^^^^
+
+handle_bad_irq is used for spurious interrupts which have no real
+handler assigned..
+
+Quirks and optimizations
+~~~~~~~~~~~~~~~~~~~~~~~~
+
+The generic functions are intended for 'clean' architectures and chips,
+which have no platform-specific IRQ handling quirks. If an architecture
+needs to implement quirks on the 'flow' level then it can do so by
+overriding the high-level irq-flow handler.
+
+Delayed interrupt disable
+~~~~~~~~~~~~~~~~~~~~~~~~~
+
+This per interrupt selectable feature, which was introduced by Russell
+King in the ARM interrupt implementation, does not mask an interrupt at
+the hardware level when disable_irq() is called. The interrupt is kept
+enabled and is masked in the flow handler when an interrupt event
+happens. This prevents losing edge interrupts on hardware which does not
+store an edge interrupt event while the interrupt is disabled at the
+hardware level. When an interrupt arrives while the IRQ_DISABLED flag
+is set, then the interrupt is masked at the hardware level and the
+IRQ_PENDING bit is set. When the interrupt is re-enabled by
+enable_irq() the pending bit is checked and if it is set, the interrupt
+is resent either via hardware or by a software resend mechanism. (It's
+necessary to enable CONFIG_HARDIRQS_SW_RESEND when you want to use
+the delayed interrupt disable feature and your hardware is not capable
+of retriggering an interrupt.) The delayed interrupt disable is not
+configurable.
+
+Chip-level hardware encapsulation
+---------------------------------
+
+The chip-level hardware descriptor structure :c:type:`irq_chip` contains all
+the direct chip relevant functions, which can be utilized by the irq flow
+implementations.
+
+-  ``irq_ack``
+
+-  ``irq_mask_ack`` - Optional, recommended for performance
+
+-  ``irq_mask``
+
+-  ``irq_unmask``
+
+-  ``irq_eoi`` - Optional, required for EOI flow handlers
+
+-  ``irq_retrigger`` - Optional
+
+-  ``irq_set_type`` - Optional
+
+-  ``irq_set_wake`` - Optional
+
+These primitives are strictly intended to mean what they say: ack means
+ACK, masking means masking of an IRQ line, etc. It is up to the flow
+handler(s) to use these basic units of low-level functionality.
+
+__do_IRQ entry point
+====================
+
+The original implementation __do_IRQ() was an alternative entry point
+for all types of interrupts. It no longer exists.
+
+This handler turned out to be not suitable for all interrupt hardware
+and was therefore reimplemented with split functionality for
+edge/level/simple/percpu interrupts. This is not only a functional
+optimization. It also shortens code paths for interrupts.
+
+Locking on SMP
+==============
+
+The locking of chip registers is up to the architecture that defines the
+chip primitives. The per-irq structure is protected via desc->lock, by
+the generic layer.
+
+Generic interrupt chip
+======================
+
+To avoid copies of identical implementations of IRQ chips the core
+provides a configurable generic interrupt chip implementation.
+Developers should check carefully whether the generic chip fits their
+needs before implementing the same functionality slightly differently
+themselves.
+
+.. kernel-doc:: kernel/irq/generic-chip.c
+   :export:
+
+Structures
+==========
+
+This chapter contains the autogenerated documentation of the structures
+which are used in the generic IRQ layer.
+
+.. kernel-doc:: include/linux/irq.h
+   :internal:
+
+.. kernel-doc:: include/linux/interrupt.h
+   :internal:
+
+Public Functions Provided
+=========================
+
+This chapter contains the autogenerated documentation of the kernel API
+functions which are exported.
+
+.. kernel-doc:: kernel/irq/manage.c
+
+.. kernel-doc:: kernel/irq/chip.c
+   :export:
+
+Internal Functions Provided
+===========================
+
+This chapter contains the autogenerated documentation of the internal
+functions.
+
+.. kernel-doc:: kernel/irq/irqdesc.c
+
+.. kernel-doc:: kernel/irq/handle.c
+
+.. kernel-doc:: kernel/irq/chip.c
+   :internal:
+
+Credits
+=======
+
+The following people have contributed to this document:
+
+1. Thomas Gleixner tglx@linutronix.de
+
+2. Ingo Molnar mingo@elte.hu
diff --git a/Documentation/core-api/gfp_mask-from-fs-io.rst b/Documentation/core-api/gfp_mask-from-fs-io.rst
new file mode 100644
index 000000000..e7c32a8de
--- /dev/null
+++ b/Documentation/core-api/gfp_mask-from-fs-io.rst
@@ -0,0 +1,68 @@
+.. _gfp_mask_from_fs_io:
+
+=================================
+GFP masks used from FS/IO context
+=================================
+
+:Date: May, 2018
+:Author: Michal Hocko <mhocko@kernel.org>
+
+Introduction
+============
+
+Code paths in the filesystem and IO stacks must be careful when
+allocating memory to prevent recursion deadlocks caused by direct
+memory reclaim calling back into the FS or IO paths and blocking on
+already held resources (e.g. locks - most commonly those used for the
+transaction context).
+
+The traditional way to avoid this deadlock problem is to clear __GFP_FS
+respectively __GFP_IO (note the latter implies clearing the first as well) in
+the gfp mask when calling an allocator. GFP_NOFS respectively GFP_NOIO can be
+used as shortcut. It turned out though that above approach has led to
+abuses when the restricted gfp mask is used "just in case" without a
+deeper consideration which leads to problems because an excessive use
+of GFP_NOFS/GFP_NOIO can lead to memory over-reclaim or other memory
+reclaim issues.
+
+New API
+========
+
+Since 4.12 we do have a generic scope API for both NOFS and NOIO context
+``memalloc_nofs_save``, ``memalloc_nofs_restore`` respectively ``memalloc_noio_save``,
+``memalloc_noio_restore`` which allow to mark a scope to be a critical
+section from a filesystem or I/O point of view. Any allocation from that
+scope will inherently drop __GFP_FS respectively __GFP_IO from the given
+mask so no memory allocation can recurse back in the FS/IO.
+
+.. kernel-doc:: include/linux/sched/mm.h
+   :functions: memalloc_nofs_save memalloc_nofs_restore
+.. kernel-doc:: include/linux/sched/mm.h
+   :functions: memalloc_noio_save memalloc_noio_restore
+
+FS/IO code then simply calls the appropriate save function before
+any critical section with respect to the reclaim is started - e.g.
+lock shared with the reclaim context or when a transaction context
+nesting would be possible via reclaim. The restore function should be
+called when the critical section ends. All that ideally along with an
+explanation what is the reclaim context for easier maintenance.
+
+Please note that the proper pairing of save/restore functions
+allows nesting so it is safe to call ``memalloc_noio_save`` or
+``memalloc_noio_restore`` respectively from an existing NOIO or NOFS
+scope.
+
+What about __vmalloc(GFP_NOFS)
+==============================
+
+vmalloc doesn't support GFP_NOFS semantic because there are hardcoded
+GFP_KERNEL allocations deep inside the allocator which are quite non-trivial
+to fix up. That means that calling ``vmalloc`` with GFP_NOFS/GFP_NOIO is
+almost always a bug. The good news is that the NOFS/NOIO semantic can be
+achieved by the scope API.
+
+In the ideal world, upper layers should already mark dangerous contexts
+and so no special care is required and vmalloc should be called without
+any problems. Sometimes if the context is not really clear or there are
+layering violations then the recommended way around that is to wrap ``vmalloc``
+by the scope API with a comment explaining the problem.
diff --git a/Documentation/core-api/idr.rst b/Documentation/core-api/idr.rst
new file mode 100644
index 000000000..2eb5afdb9
--- /dev/null
+++ b/Documentation/core-api/idr.rst
@@ -0,0 +1,81 @@
+.. SPDX-License-Identifier: GPL-2.0+
+
+=============
+ID Allocation
+=============
+
+:Author: Matthew Wilcox
+
+Overview
+========
+
+A common problem to solve is allocating identifiers (IDs); generally
+small numbers which identify a thing.  Examples include file descriptors,
+process IDs, packet identifiers in networking protocols, SCSI tags
+and device instance numbers.  The IDR and the IDA provide a reasonable
+solution to the problem to avoid everybody inventing their own.  The IDR
+provides the ability to map an ID to a pointer, while the IDA provides
+only ID allocation, and as a result is much more memory-efficient.
+
+IDR usage
+=========
+
+Start by initialising an IDR, either with DEFINE_IDR()
+for statically allocated IDRs or idr_init() for dynamically
+allocated IDRs.
+
+You can call idr_alloc() to allocate an unused ID.  Look up
+the pointer you associated with the ID by calling idr_find()
+and free the ID by calling idr_remove().
+
+If you need to change the pointer associated with an ID, you can call
+idr_replace().  One common reason to do this is to reserve an
+ID by passing a ``NULL`` pointer to the allocation function; initialise the
+object with the reserved ID and finally insert the initialised object
+into the IDR.
+
+Some users need to allocate IDs larger than ``INT_MAX``.  So far all of
+these users have been content with a ``UINT_MAX`` limit, and they use
+idr_alloc_u32().  If you need IDs that will not fit in a u32,
+we will work with you to address your needs.
+
+If you need to allocate IDs sequentially, you can use
+idr_alloc_cyclic().  The IDR becomes less efficient when dealing
+with larger IDs, so using this function comes at a slight cost.
+
+To perform an action on all pointers used by the IDR, you can
+either use the callback-based idr_for_each() or the
+iterator-style idr_for_each_entry().  You may need to use
+idr_for_each_entry_continue() to continue an iteration.  You can
+also use idr_get_next() if the iterator doesn't fit your needs.
+
+When you have finished using an IDR, you can call idr_destroy()
+to release the memory used by the IDR.  This will not free the objects
+pointed to from the IDR; if you want to do that, use one of the iterators
+to do it.
+
+You can use idr_is_empty() to find out whether there are any
+IDs currently allocated.
+
+If you need to take a lock while allocating a new ID from the IDR,
+you may need to pass a restrictive set of GFP flags, which can lead
+to the IDR being unable to allocate memory.  To work around this,
+you can call idr_preload() before taking the lock, and then
+idr_preload_end() after the allocation.
+
+.. kernel-doc:: include/linux/idr.h
+   :doc: idr sync
+
+IDA usage
+=========
+
+.. kernel-doc:: lib/idr.c
+   :doc: IDA description
+
+Functions and structures
+========================
+
+.. kernel-doc:: include/linux/idr.h
+   :functions:
+.. kernel-doc:: lib/idr.c
+   :functions:
diff --git a/Documentation/core-api/index.rst b/Documentation/core-api/index.rst
new file mode 100644
index 000000000..69171b179
--- /dev/null
+++ b/Documentation/core-api/index.rst
@@ -0,0 +1,123 @@
+======================
+Core API Documentation
+======================
+
+This is the beginning of a manual for core kernel APIs.  The conversion
+(and writing!) of documents for this manual is much appreciated!
+
+Core utilities
+==============
+
+This section has general and "core core" documentation.  The first is a
+massive grab-bag of kerneldoc info left over from the docbook days; it
+should really be broken up someday when somebody finds the energy to do
+it.
+
+.. toctree::
+   :maxdepth: 1
+
+   kernel-api
+   workqueue
+   printk-basics
+   printk-formats
+   symbol-namespaces
+
+Data structures and low-level utilities
+=======================================
+
+Library functionality that is used throughout the kernel.
+
+.. toctree::
+   :maxdepth: 1
+
+   kobject
+   kref
+   assoc_array
+   xarray
+   idr
+   circular-buffers
+   rbtree
+   generic-radix-tree
+   packing
+   bus-virt-phys-mapping
+   this_cpu_ops
+   timekeeping
+   errseq
+
+Concurrency primitives
+======================
+
+How Linux keeps everything from happening at the same time.  See
+:doc:`/locking/index` for more related documentation.
+
+.. toctree::
+   :maxdepth: 1
+
+   atomic_ops
+   refcount-vs-atomic
+   irq/index
+   local_ops
+   padata
+   ../RCU/index
+
+Low-level hardware management
+=============================
+
+Cache management, managing CPU hotplug, etc.
+
+.. toctree::
+   :maxdepth: 1
+
+   cachetlb
+   cpu_hotplug
+   memory-hotplug
+   genericirq
+   protection-keys
+
+Memory management
+=================
+
+How to allocate and use memory in the kernel.  Note that there is a lot
+more memory-management documentation in :doc:`/vm/index`.
+
+.. toctree::
+   :maxdepth: 1
+
+   memory-allocation
+   unaligned-memory-access
+   dma-api
+   dma-api-howto
+   dma-attributes
+   dma-isa-lpc
+   mm-api
+   genalloc
+   pin_user_pages
+   boot-time-mm
+   gfp_mask-from-fs-io
+
+Interfaces for kernel debugging
+===============================
+
+.. toctree::
+   :maxdepth: 1
+
+   debug-objects
+   tracepoint
+   debugging-via-ohci1394
+
+Everything else
+===============
+
+Documents that don't fit elsewhere or which have yet to be categorized.
+
+.. toctree::
+   :maxdepth: 1
+
+   librs
+
+.. only:: subproject and html
+
+   Indices
+   =======
+
+   * :ref:`genindex`
diff --git a/Documentation/core-api/irq/concepts.rst b/Documentation/core-api/irq/concepts.rst
new file mode 100644
index 000000000..4273806a6
--- /dev/null
+++ b/Documentation/core-api/irq/concepts.rst
@@ -0,0 +1,24 @@
+===============
+What is an IRQ?
+===============
+
+An IRQ is an interrupt request from a device.
+Currently they can come in over a pin, or over a packet.
+Several devices may be connected to the same pin thus
+sharing an IRQ.
+
+An IRQ number is a kernel identifier used to talk about a hardware
+interrupt source.  Typically this is an index into the global irq_desc
+array, but except for what linux/interrupt.h implements the details
+are architecture specific.
+
+An IRQ number is an enumeration of the possible interrupt sources on a
+machine.  Typically what is enumerated is the number of input pins on
+all of the interrupt controller in the system.  In the case of ISA
+what is enumerated are the 16 input pins on the two i8259 interrupt
+controllers.
+
+Architectures can assign additional meaning to the IRQ numbers, and
+are encouraged to in the case  where there is any manual configuration
+of the hardware involved.  The ISA IRQs are a classic example of
+assigning this kind of additional meaning.
diff --git a/Documentation/core-api/irq/index.rst b/Documentation/core-api/irq/index.rst
new file mode 100644
index 000000000..0d65d11e5
--- /dev/null
+++ b/Documentation/core-api/irq/index.rst
@@ -0,0 +1,11 @@
+====
+IRQs
+====
+
+.. toctree::
+   :maxdepth: 1
+
+   concepts
+   irq-affinity
+   irq-domain
+   irqflags-tracing
diff --git a/Documentation/core-api/irq/irq-affinity.rst b/Documentation/core-api/irq/irq-affinity.rst
new file mode 100644
index 000000000..29da50008
--- /dev/null
+++ b/Documentation/core-api/irq/irq-affinity.rst
@@ -0,0 +1,70 @@
+================
+SMP IRQ affinity
+================
+
+ChangeLog:
+	- Started by Ingo Molnar <mingo@redhat.com>
+	- Update by Max Krasnyansky <maxk@qualcomm.com>
+
+
+/proc/irq/IRQ#/smp_affinity and /proc/irq/IRQ#/smp_affinity_list specify
+which target CPUs are permitted for a given IRQ source.  It's a bitmask
+(smp_affinity) or cpu list (smp_affinity_list) of allowed CPUs.  It's not
+allowed to turn off all CPUs, and if an IRQ controller does not support
+IRQ affinity then the value will not change from the default of all cpus.
+
+/proc/irq/default_smp_affinity specifies default affinity mask that applies
+to all non-active IRQs. Once IRQ is allocated/activated its affinity bitmask
+will be set to the default mask. It can then be changed as described above.
+Default mask is 0xffffffff.
+
+Here is an example of restricting IRQ44 (eth1) to CPU0-3 then restricting
+it to CPU4-7 (this is an 8-CPU SMP box)::
+
+	[root@moon 44]# cd /proc/irq/44
+	[root@moon 44]# cat smp_affinity
+	ffffffff
+
+	[root@moon 44]# echo 0f > smp_affinity
+	[root@moon 44]# cat smp_affinity
+	0000000f
+	[root@moon 44]# ping -f h
+	PING hell (195.4.7.3): 56 data bytes
+	...
+	--- hell ping statistics ---
+	6029 packets transmitted, 6027 packets received, 0% packet loss
+	round-trip min/avg/max = 0.1/0.1/0.4 ms
+	[root@moon 44]# cat /proc/interrupts | grep 'CPU\|44:'
+		CPU0       CPU1       CPU2       CPU3      CPU4       CPU5        CPU6       CPU7
+	44:       1068       1785       1785       1783         0          0           0         0    IO-APIC-level  eth1
+
+As can be seen from the line above IRQ44 was delivered only to the first four
+processors (0-3).
+Now lets restrict that IRQ to CPU(4-7).
+
+::
+
+	[root@moon 44]# echo f0 > smp_affinity
+	[root@moon 44]# cat smp_affinity
+	000000f0
+	[root@moon 44]# ping -f h
+	PING hell (195.4.7.3): 56 data bytes
+	..
+	--- hell ping statistics ---
+	2779 packets transmitted, 2777 packets received, 0% packet loss
+	round-trip min/avg/max = 0.1/0.5/585.4 ms
+	[root@moon 44]# cat /proc/interrupts |  'CPU\|44:'
+		CPU0       CPU1       CPU2       CPU3      CPU4       CPU5        CPU6       CPU7
+	44:       1068       1785       1785       1783      1784       1069        1070       1069   IO-APIC-level  eth1
+
+This time around IRQ44 was delivered only to the last four processors.
+i.e counters for the CPU0-3 did not change.
+
+Here is an example of limiting that same irq (44) to cpus 1024 to 1031::
+
+	[root@moon 44]# echo 1024-1031 > smp_affinity_list
+	[root@moon 44]# cat smp_affinity_list
+	1024-1031
+
+Note that to do this with a bitmask would require 32 bitmasks of zero
+to follow the pertinent one.
diff --git a/Documentation/core-api/irq/irq-domain.rst b/Documentation/core-api/irq/irq-domain.rst
new file mode 100644
index 000000000..096db12f3
--- /dev/null
+++ b/Documentation/core-api/irq/irq-domain.rst
@@ -0,0 +1,270 @@
+===============================================
+The irq_domain interrupt number mapping library
+===============================================
+
+The current design of the Linux kernel uses a single large number
+space where each separate IRQ source is assigned a different number.
+This is simple when there is only one interrupt controller, but in
+systems with multiple interrupt controllers the kernel must ensure
+that each one gets assigned non-overlapping allocations of Linux
+IRQ numbers.
+
+The number of interrupt controllers registered as unique irqchips
+show a rising tendency: for example subdrivers of different kinds
+such as GPIO controllers avoid reimplementing identical callback
+mechanisms as the IRQ core system by modelling their interrupt
+handlers as irqchips, i.e. in effect cascading interrupt controllers.
+
+Here the interrupt number loose all kind of correspondence to
+hardware interrupt numbers: whereas in the past, IRQ numbers could
+be chosen so they matched the hardware IRQ line into the root
+interrupt controller (i.e. the component actually fireing the
+interrupt line to the CPU) nowadays this number is just a number.
+
+For this reason we need a mechanism to separate controller-local
+interrupt numbers, called hardware irq's, from Linux IRQ numbers.
+
+The irq_alloc_desc*() and irq_free_desc*() APIs provide allocation of
+irq numbers, but they don't provide any support for reverse mapping of
+the controller-local IRQ (hwirq) number into the Linux IRQ number
+space.
+
+The irq_domain library adds mapping between hwirq and IRQ numbers on
+top of the irq_alloc_desc*() API.  An irq_domain to manage mapping is
+preferred over interrupt controller drivers open coding their own
+reverse mapping scheme.
+
+irq_domain also implements translation from an abstract irq_fwspec
+structure to hwirq numbers (Device Tree and ACPI GSI so far), and can
+be easily extended to support other IRQ topology data sources.
+
+irq_domain usage
+================
+
+An interrupt controller driver creates and registers an irq_domain by
+calling one of the irq_domain_add_*() functions (each mapping method
+has a different allocator function, more on that later).  The function
+will return a pointer to the irq_domain on success.  The caller must
+provide the allocator function with an irq_domain_ops structure.
+
+In most cases, the irq_domain will begin empty without any mappings
+between hwirq and IRQ numbers.  Mappings are added to the irq_domain
+by calling irq_create_mapping() which accepts the irq_domain and a
+hwirq number as arguments.  If a mapping for the hwirq doesn't already
+exist then it will allocate a new Linux irq_desc, associate it with
+the hwirq, and call the .map() callback so the driver can perform any
+required hardware setup.
+
+When an interrupt is received, irq_find_mapping() function should
+be used to find the Linux IRQ number from the hwirq number.
+
+The irq_create_mapping() function must be called *atleast once*
+before any call to irq_find_mapping(), lest the descriptor will not
+be allocated.
+
+If the driver has the Linux IRQ number or the irq_data pointer, and
+needs to know the associated hwirq number (such as in the irq_chip
+callbacks) then it can be directly obtained from irq_data->hwirq.
+
+Types of irq_domain mappings
+============================
+
+There are several mechanisms available for reverse mapping from hwirq
+to Linux irq, and each mechanism uses a different allocation function.
+Which reverse map type should be used depends on the use case.  Each
+of the reverse map types are described below:
+
+Linear
+------
+
+::
+
+	irq_domain_add_linear()
+	irq_domain_create_linear()
+
+The linear reverse map maintains a fixed size table indexed by the
+hwirq number.  When a hwirq is mapped, an irq_desc is allocated for
+the hwirq, and the IRQ number is stored in the table.
+
+The Linear map is a good choice when the maximum number of hwirqs is
+fixed and a relatively small number (~ < 256).  The advantages of this
+map are fixed time lookup for IRQ numbers, and irq_descs are only
+allocated for in-use IRQs.  The disadvantage is that the table must be
+as large as the largest possible hwirq number.
+
+irq_domain_add_linear() and irq_domain_create_linear() are functionally
+equivalent, except for the first argument is different - the former
+accepts an Open Firmware specific 'struct device_node', while the latter
+accepts a more general abstraction 'struct fwnode_handle'.
+
+The majority of drivers should use the linear map.
+
+Tree
+----
+
+::
+
+	irq_domain_add_tree()
+	irq_domain_create_tree()
+
+The irq_domain maintains a radix tree map from hwirq numbers to Linux
+IRQs.  When an hwirq is mapped, an irq_desc is allocated and the
+hwirq is used as the lookup key for the radix tree.
+
+The tree map is a good choice if the hwirq number can be very large
+since it doesn't need to allocate a table as large as the largest
+hwirq number.  The disadvantage is that hwirq to IRQ number lookup is
+dependent on how many entries are in the table.
+
+irq_domain_add_tree() and irq_domain_create_tree() are functionally
+equivalent, except for the first argument is different - the former
+accepts an Open Firmware specific 'struct device_node', while the latter
+accepts a more general abstraction 'struct fwnode_handle'.
+
+Very few drivers should need this mapping.
+
+No Map
+------
+
+::
+
+	irq_domain_add_nomap()
+
+The No Map mapping is to be used when the hwirq number is
+programmable in the hardware.  In this case it is best to program the
+Linux IRQ number into the hardware itself so that no mapping is
+required.  Calling irq_create_direct_mapping() will allocate a Linux
+IRQ number and call the .map() callback so that driver can program the
+Linux IRQ number into the hardware.
+
+Most drivers cannot use this mapping.
+
+Legacy
+------
+
+::
+
+	irq_domain_add_simple()
+	irq_domain_add_legacy()
+	irq_domain_add_legacy_isa()
+
+The Legacy mapping is a special case for drivers that already have a
+range of irq_descs allocated for the hwirqs.  It is used when the
+driver cannot be immediately converted to use the linear mapping.  For
+example, many embedded system board support files use a set of #defines
+for IRQ numbers that are passed to struct device registrations.  In that
+case the Linux IRQ numbers cannot be dynamically assigned and the legacy
+mapping should be used.
+
+The legacy map assumes a contiguous range of IRQ numbers has already
+been allocated for the controller and that the IRQ number can be
+calculated by adding a fixed offset to the hwirq number, and
+visa-versa.  The disadvantage is that it requires the interrupt
+controller to manage IRQ allocations and it requires an irq_desc to be
+allocated for every hwirq, even if it is unused.
+
+The legacy map should only be used if fixed IRQ mappings must be
+supported.  For example, ISA controllers would use the legacy map for
+mapping Linux IRQs 0-15 so that existing ISA drivers get the correct IRQ
+numbers.
+
+Most users of legacy mappings should use irq_domain_add_simple() which
+will use a legacy domain only if an IRQ range is supplied by the
+system and will otherwise use a linear domain mapping. The semantics
+of this call are such that if an IRQ range is specified then
+descriptors will be allocated on-the-fly for it, and if no range is
+specified it will fall through to irq_domain_add_linear() which means
+*no* irq descriptors will be allocated.
+
+A typical use case for simple domains is where an irqchip provider
+is supporting both dynamic and static IRQ assignments.
+
+In order to avoid ending up in a situation where a linear domain is
+used and no descriptor gets allocated it is very important to make sure
+that the driver using the simple domain call irq_create_mapping()
+before any irq_find_mapping() since the latter will actually work
+for the static IRQ assignment case.
+
+Hierarchy IRQ domain
+--------------------
+
+On some architectures, there may be multiple interrupt controllers
+involved in delivering an interrupt from the device to the target CPU.
+Let's look at a typical interrupt delivering path on x86 platforms::
+
+  Device --> IOAPIC -> Interrupt remapping Controller -> Local APIC -> CPU
+
+There are three interrupt controllers involved:
+
+1) IOAPIC controller
+2) Interrupt remapping controller
+3) Local APIC controller
+
+To support such a hardware topology and make software architecture match
+hardware architecture, an irq_domain data structure is built for each
+interrupt controller and those irq_domains are organized into hierarchy.
+When building irq_domain hierarchy, the irq_domain near to the device is
+child and the irq_domain near to CPU is parent. So a hierarchy structure
+as below will be built for the example above::
+
+	CPU Vector irq_domain (root irq_domain to manage CPU vectors)
+		^
+		|
+	Interrupt Remapping irq_domain (manage irq_remapping entries)
+		^
+		|
+	IOAPIC irq_domain (manage IOAPIC delivery entries/pins)
+
+There are four major interfaces to use hierarchy irq_domain:
+
+1) irq_domain_alloc_irqs(): allocate IRQ descriptors and interrupt
+   controller related resources to deliver these interrupts.
+2) irq_domain_free_irqs(): free IRQ descriptors and interrupt controller
+   related resources associated with these interrupts.
+3) irq_domain_activate_irq(): activate interrupt controller hardware to
+   deliver the interrupt.
+4) irq_domain_deactivate_irq(): deactivate interrupt controller hardware
+   to stop delivering the interrupt.
+
+Following changes are needed to support hierarchy irq_domain:
+
+1) a new field 'parent' is added to struct irq_domain; it's used to
+   maintain irq_domain hierarchy information.
+2) a new field 'parent_data' is added to struct irq_data; it's used to
+   build hierarchy irq_data to match hierarchy irq_domains. The irq_data
+   is used to store irq_domain pointer and hardware irq number.
+3) new callbacks are added to struct irq_domain_ops to support hierarchy
+   irq_domain operations.
+
+With support of hierarchy irq_domain and hierarchy irq_data ready, an
+irq_domain structure is built for each interrupt controller, and an
+irq_data structure is allocated for each irq_domain associated with an
+IRQ. Now we could go one step further to support stacked(hierarchy)
+irq_chip. That is, an irq_chip is associated with each irq_data along
+the hierarchy. A child irq_chip may implement a required action by
+itself or by cooperating with its parent irq_chip.
+
+With stacked irq_chip, interrupt controller driver only needs to deal
+with the hardware managed by itself and may ask for services from its
+parent irq_chip when needed. So we could achieve a much cleaner
+software architecture.
+
+For an interrupt controller driver to support hierarchy irq_domain, it
+needs to:
+
+1) Implement irq_domain_ops.alloc and irq_domain_ops.free
+2) Optionally implement irq_domain_ops.activate and
+   irq_domain_ops.deactivate.
+3) Optionally implement an irq_chip to manage the interrupt controller
+   hardware.
+4) No need to implement irq_domain_ops.map and irq_domain_ops.unmap,
+   they are unused with hierarchy irq_domain.
+
+Hierarchy irq_domain is in no way x86 specific, and is heavily used to
+support other architectures, such as ARM, ARM64 etc.
+
+Debugging
+=========
+
+Most of the internals of the IRQ subsystem are exposed in debugfs by
+turning CONFIG_GENERIC_IRQ_DEBUGFS on.
diff --git a/Documentation/core-api/irq/irqflags-tracing.rst b/Documentation/core-api/irq/irqflags-tracing.rst
new file mode 100644
index 000000000..bdd208259
--- /dev/null
+++ b/Documentation/core-api/irq/irqflags-tracing.rst
@@ -0,0 +1,52 @@
+=======================
+IRQ-flags state tracing
+=======================
+
+:Author: started by Ingo Molnar <mingo@redhat.com>
+
+The "irq-flags tracing" feature "traces" hardirq and softirq state, in
+that it gives interested subsystems an opportunity to be notified of
+every hardirqs-off/hardirqs-on, softirqs-off/softirqs-on event that
+happens in the kernel.
+
+CONFIG_TRACE_IRQFLAGS_SUPPORT is needed for CONFIG_PROVE_SPIN_LOCKING
+and CONFIG_PROVE_RW_LOCKING to be offered by the generic lock debugging
+code. Otherwise only CONFIG_PROVE_MUTEX_LOCKING and
+CONFIG_PROVE_RWSEM_LOCKING will be offered on an architecture - these
+are locking APIs that are not used in IRQ context. (the one exception
+for rwsems is worked around)
+
+Architecture support for this is certainly not in the "trivial"
+category, because lots of lowlevel assembly code deal with irq-flags
+state changes. But an architecture can be irq-flags-tracing enabled in a
+rather straightforward and risk-free manner.
+
+Architectures that want to support this need to do a couple of
+code-organizational changes first:
+
+- add and enable TRACE_IRQFLAGS_SUPPORT in their arch level Kconfig file
+
+and then a couple of functional changes are needed as well to implement
+irq-flags-tracing support:
+
+- in lowlevel entry code add (build-conditional) calls to the
+  trace_hardirqs_off()/trace_hardirqs_on() functions. The lock validator
+  closely guards whether the 'real' irq-flags matches the 'virtual'
+  irq-flags state, and complains loudly (and turns itself off) if the
+  two do not match. Usually most of the time for arch support for
+  irq-flags-tracing is spent in this state: look at the lockdep
+  complaint, try to figure out the assembly code we did not cover yet,
+  fix and repeat. Once the system has booted up and works without a
+  lockdep complaint in the irq-flags-tracing functions arch support is
+  complete.
+- if the architecture has non-maskable interrupts then those need to be
+  excluded from the irq-tracing [and lock validation] mechanism via
+  lockdep_off()/lockdep_on().
+
+In general there is no risk from having an incomplete irq-flags-tracing
+implementation in an architecture: lockdep will detect that and will
+turn itself off. I.e. the lock validator will still be reliable. There
+should be no crashes due to irq-tracing bugs. (except if the assembly
+changes break other code by modifying conditions or registers that
+shouldn't be)
+
diff --git a/Documentation/core-api/kernel-api.rst b/Documentation/core-api/kernel-api.rst
new file mode 100644
index 000000000..2a7444e3a
--- /dev/null
+++ b/Documentation/core-api/kernel-api.rst
@@ -0,0 +1,389 @@
+====================
+The Linux Kernel API
+====================
+
+
+List Management Functions
+=========================
+
+.. kernel-doc:: include/linux/list.h
+   :internal:
+
+Basic C Library Functions
+=========================
+
+When writing drivers, you cannot in general use routines which are from
+the C Library. Some of the functions have been found generally useful
+and they are listed below. The behaviour of these functions may vary
+slightly from those defined by ANSI, and these deviations are noted in
+the text.
+
+String Conversions
+------------------
+
+.. kernel-doc:: lib/vsprintf.c
+   :export:
+
+.. kernel-doc:: include/linux/kstrtox.h
+   :functions: kstrtol kstrtoul
+
+.. kernel-doc:: lib/kstrtox.c
+   :export:
+
+.. kernel-doc:: lib/string_helpers.c
+   :export:
+
+String Manipulation
+-------------------
+
+.. kernel-doc:: lib/string.c
+   :export:
+
+.. kernel-doc:: include/linux/string.h
+   :internal:
+
+.. kernel-doc:: mm/util.c
+   :functions: kstrdup kstrdup_const kstrndup kmemdup kmemdup_nul memdup_user
+               vmemdup_user strndup_user memdup_user_nul
+
+Basic Kernel Library Functions
+==============================
+
+The Linux kernel provides more basic utility functions.
+
+Bit Operations
+--------------
+
+.. kernel-doc:: include/asm-generic/bitops/instrumented-atomic.h
+   :internal:
+
+.. kernel-doc:: include/asm-generic/bitops/instrumented-non-atomic.h
+   :internal:
+
+.. kernel-doc:: include/asm-generic/bitops/instrumented-lock.h
+   :internal:
+
+Bitmap Operations
+-----------------
+
+.. kernel-doc:: lib/bitmap.c
+   :doc: bitmap introduction
+
+.. kernel-doc:: include/linux/bitmap.h
+   :doc: declare bitmap
+
+.. kernel-doc:: include/linux/bitmap.h
+   :doc: bitmap overview
+
+.. kernel-doc:: include/linux/bitmap.h
+   :doc: bitmap bitops
+
+.. kernel-doc:: lib/bitmap.c
+   :export:
+
+.. kernel-doc:: lib/bitmap.c
+   :internal:
+
+.. kernel-doc:: include/linux/bitmap.h
+   :internal:
+
+Command-line Parsing
+--------------------
+
+.. kernel-doc:: lib/cmdline.c
+   :export:
+
+Sorting
+-------
+
+.. kernel-doc:: lib/sort.c
+   :export:
+
+.. kernel-doc:: lib/list_sort.c
+   :export:
+
+Text Searching
+--------------
+
+.. kernel-doc:: lib/textsearch.c
+   :doc: ts_intro
+
+.. kernel-doc:: lib/textsearch.c
+   :export:
+
+.. kernel-doc:: include/linux/textsearch.h
+   :functions: textsearch_find textsearch_next \
+               textsearch_get_pattern textsearch_get_pattern_len
+
+CRC and Math Functions in Linux
+===============================
+
+CRC Functions
+-------------
+
+.. kernel-doc:: lib/crc4.c
+   :export:
+
+.. kernel-doc:: lib/crc7.c
+   :export:
+
+.. kernel-doc:: lib/crc8.c
+   :export:
+
+.. kernel-doc:: lib/crc16.c
+   :export:
+
+.. kernel-doc:: lib/crc32.c
+
+.. kernel-doc:: lib/crc-ccitt.c
+   :export:
+
+.. kernel-doc:: lib/crc-itu-t.c
+   :export:
+
+Base 2 log and power Functions
+------------------------------
+
+.. kernel-doc:: include/linux/log2.h
+   :internal:
+
+Integer power Functions
+-----------------------
+
+.. kernel-doc:: lib/math/int_pow.c
+   :export:
+
+.. kernel-doc:: lib/math/int_sqrt.c
+   :export:
+
+Division Functions
+------------------
+
+.. kernel-doc:: include/asm-generic/div64.h
+   :functions: do_div
+
+.. kernel-doc:: include/linux/math64.h
+   :internal:
+
+.. kernel-doc:: lib/math/div64.c
+   :functions: div_s64_rem div64_u64_rem div64_u64 div64_s64
+
+.. kernel-doc:: lib/math/gcd.c
+   :export:
+
+UUID/GUID
+---------
+
+.. kernel-doc:: lib/uuid.c
+   :export:
+
+Kernel IPC facilities
+=====================
+
+IPC utilities
+-------------
+
+.. kernel-doc:: ipc/util.c
+   :internal:
+
+FIFO Buffer
+===========
+
+kfifo interface
+---------------
+
+.. kernel-doc:: include/linux/kfifo.h
+   :internal:
+
+relay interface support
+=======================
+
+Relay interface support is designed to provide an efficient mechanism
+for tools and facilities to relay large amounts of data from kernel
+space to user space.
+
+relay interface
+---------------
+
+.. kernel-doc:: kernel/relay.c
+   :export:
+
+.. kernel-doc:: kernel/relay.c
+   :internal:
+
+Module Support
+==============
+
+Module Loading
+--------------
+
+.. kernel-doc:: kernel/kmod.c
+   :export:
+
+Inter Module support
+--------------------
+
+Refer to the file kernel/module.c for more information.
+
+Hardware Interfaces
+===================
+
+DMA Channels
+------------
+
+.. kernel-doc:: kernel/dma.c
+   :export:
+
+Resources Management
+--------------------
+
+.. kernel-doc:: kernel/resource.c
+   :internal:
+
+.. kernel-doc:: kernel/resource.c
+   :export:
+
+MTRR Handling
+-------------
+
+.. kernel-doc:: arch/x86/kernel/cpu/mtrr/mtrr.c
+   :export:
+
+Security Framework
+==================
+
+.. kernel-doc:: security/security.c
+   :internal:
+
+.. kernel-doc:: security/inode.c
+   :export:
+
+Audit Interfaces
+================
+
+.. kernel-doc:: kernel/audit.c
+   :export:
+
+.. kernel-doc:: kernel/auditsc.c
+   :internal:
+
+.. kernel-doc:: kernel/auditfilter.c
+   :internal:
+
+Accounting Framework
+====================
+
+.. kernel-doc:: kernel/acct.c
+   :internal:
+
+Block Devices
+=============
+
+.. kernel-doc:: block/blk-core.c
+   :export:
+
+.. kernel-doc:: block/blk-core.c
+   :internal:
+
+.. kernel-doc:: block/blk-map.c
+   :export:
+
+.. kernel-doc:: block/blk-sysfs.c
+   :internal:
+
+.. kernel-doc:: block/blk-settings.c
+   :export:
+
+.. kernel-doc:: block/blk-exec.c
+   :export:
+
+.. kernel-doc:: block/blk-flush.c
+   :export:
+
+.. kernel-doc:: block/blk-lib.c
+   :export:
+
+.. kernel-doc:: block/blk-integrity.c
+   :export:
+
+.. kernel-doc:: kernel/trace/blktrace.c
+   :internal:
+
+.. kernel-doc:: block/genhd.c
+   :internal:
+
+.. kernel-doc:: block/genhd.c
+   :export:
+
+Char devices
+============
+
+.. kernel-doc:: fs/char_dev.c
+   :export:
+
+Clock Framework
+===============
+
+The clock framework defines programming interfaces to support software
+management of the system clock tree. This framework is widely used with
+System-On-Chip (SOC) platforms to support power management and various
+devices which may need custom clock rates. Note that these "clocks"
+don't relate to timekeeping or real time clocks (RTCs), each of which
+have separate frameworks. These :c:type:`struct clk <clk>`
+instances may be used to manage for example a 96 MHz signal that is used
+to shift bits into and out of peripherals or busses, or otherwise
+trigger synchronous state machine transitions in system hardware.
+
+Power management is supported by explicit software clock gating: unused
+clocks are disabled, so the system doesn't waste power changing the
+state of transistors that aren't in active use. On some systems this may
+be backed by hardware clock gating, where clocks are gated without being
+disabled in software. Sections of chips that are powered but not clocked
+may be able to retain their last state. This low power state is often
+called a *retention mode*. This mode still incurs leakage currents,
+especially with finer circuit geometries, but for CMOS circuits power is
+mostly used by clocked state changes.
+
+Power-aware drivers only enable their clocks when the device they manage
+is in active use. Also, system sleep states often differ according to
+which clock domains are active: while a "standby" state may allow wakeup
+from several active domains, a "mem" (suspend-to-RAM) state may require
+a more wholesale shutdown of clocks derived from higher speed PLLs and
+oscillators, limiting the number of possible wakeup event sources. A
+driver's suspend method may need to be aware of system-specific clock
+constraints on the target sleep state.
+
+Some platforms support programmable clock generators. These can be used
+by external chips of various kinds, such as other CPUs, multimedia
+codecs, and devices with strict requirements for interface clocking.
+
+.. kernel-doc:: include/linux/clk.h
+   :internal:
+
+Synchronization Primitives
+==========================
+
+Read-Copy Update (RCU)
+----------------------
+
+.. kernel-doc:: include/linux/rcupdate.h
+
+.. kernel-doc:: kernel/rcu/tree.c
+
+.. kernel-doc:: kernel/rcu/tree_exp.h
+
+.. kernel-doc:: kernel/rcu/update.c
+
+.. kernel-doc:: include/linux/srcu.h
+
+.. kernel-doc:: kernel/rcu/srcutree.c
+
+.. kernel-doc:: include/linux/rculist_bl.h
+
+.. kernel-doc:: include/linux/rculist.h
+
+.. kernel-doc:: include/linux/rculist_nulls.h
+
+.. kernel-doc:: include/linux/rcu_sync.h
+
+.. kernel-doc:: kernel/rcu/sync.c
diff --git a/Documentation/core-api/kobject.rst b/Documentation/core-api/kobject.rst
new file mode 100644
index 000000000..2739f8b72
--- /dev/null
+++ b/Documentation/core-api/kobject.rst
@@ -0,0 +1,436 @@
+=====================================================================
+Everything you never wanted to know about kobjects, ksets, and ktypes
+=====================================================================
+
+:Author: Greg Kroah-Hartman <gregkh@linuxfoundation.org>
+:Last updated: December 19, 2007
+
+Based on an original article by Jon Corbet for lwn.net written October 1,
+2003 and located at https://lwn.net/Articles/51437/
+
+Part of the difficulty in understanding the driver model - and the kobject
+abstraction upon which it is built - is that there is no obvious starting
+place. Dealing with kobjects requires understanding a few different types,
+all of which make reference to each other. In an attempt to make things
+easier, we'll take a multi-pass approach, starting with vague terms and
+adding detail as we go. To that end, here are some quick definitions of
+some terms we will be working with.
+
+ - A kobject is an object of type struct kobject.  Kobjects have a name
+   and a reference count.  A kobject also has a parent pointer (allowing
+   objects to be arranged into hierarchies), a specific type, and,
+   usually, a representation in the sysfs virtual filesystem.
+
+   Kobjects are generally not interesting on their own; instead, they are
+   usually embedded within some other structure which contains the stuff
+   the code is really interested in.
+
+   No structure should **EVER** have more than one kobject embedded within it.
+   If it does, the reference counting for the object is sure to be messed
+   up and incorrect, and your code will be buggy.  So do not do this.
+
+ - A ktype is the type of object that embeds a kobject.  Every structure
+   that embeds a kobject needs a corresponding ktype.  The ktype controls
+   what happens to the kobject when it is created and destroyed.
+
+ - A kset is a group of kobjects.  These kobjects can be of the same ktype
+   or belong to different ktypes.  The kset is the basic container type for
+   collections of kobjects. Ksets contain their own kobjects, but you can
+   safely ignore that implementation detail as the kset core code handles
+   this kobject automatically.
+
+   When you see a sysfs directory full of other directories, generally each
+   of those directories corresponds to a kobject in the same kset.
+
+We'll look at how to create and manipulate all of these types. A bottom-up
+approach will be taken, so we'll go back to kobjects.
+
+
+Embedding kobjects
+==================
+
+It is rare for kernel code to create a standalone kobject, with one major
+exception explained below.  Instead, kobjects are used to control access to
+a larger, domain-specific object.  To this end, kobjects will be found
+embedded in other structures.  If you are used to thinking of things in
+object-oriented terms, kobjects can be seen as a top-level, abstract class
+from which other classes are derived.  A kobject implements a set of
+capabilities which are not particularly useful by themselves, but are
+nice to have in other objects.  The C language does not allow for the
+direct expression of inheritance, so other techniques - such as structure
+embedding - must be used.
+
+(As an aside, for those familiar with the kernel linked list implementation,
+this is analogous as to how "list_head" structs are rarely useful on
+their own, but are invariably found embedded in the larger objects of
+interest.)
+
+So, for example, the UIO code in ``drivers/uio/uio.c`` has a structure that
+defines the memory region associated with a uio device::
+
+    struct uio_map {
+            struct kobject kobj;
+            struct uio_mem *mem;
+    };
+
+If you have a struct uio_map structure, finding its embedded kobject is
+just a matter of using the kobj member.  Code that works with kobjects will
+often have the opposite problem, however: given a struct kobject pointer,
+what is the pointer to the containing structure?  You must avoid tricks
+(such as assuming that the kobject is at the beginning of the structure)
+and, instead, use the container_of() macro, found in ``<linux/kernel.h>``::
+
+    container_of(ptr, type, member)
+
+where:
+
+  * ``ptr`` is the pointer to the embedded kobject,
+  * ``type`` is the type of the containing structure, and
+  * ``member`` is the name of the structure field to which ``pointer`` points.
+
+The return value from container_of() is a pointer to the corresponding
+container type. So, for example, a pointer ``kp`` to a struct kobject
+embedded **within** a struct uio_map could be converted to a pointer to the
+**containing** uio_map structure with::
+
+    struct uio_map *u_map = container_of(kp, struct uio_map, kobj);
+
+For convenience, programmers often define a simple macro for **back-casting**
+kobject pointers to the containing type.  Exactly this happens in the
+earlier ``drivers/uio/uio.c``, as you can see here::
+
+    struct uio_map {
+            struct kobject kobj;
+            struct uio_mem *mem;
+    };
+
+    #define to_map(map) container_of(map, struct uio_map, kobj)
+
+where the macro argument "map" is a pointer to the struct kobject in
+question.  That macro is subsequently invoked with::
+
+    struct uio_map *map = to_map(kobj);
+
+
+Initialization of kobjects
+==========================
+
+Code which creates a kobject must, of course, initialize that object. Some
+of the internal fields are setup with a (mandatory) call to kobject_init()::
+
+    void kobject_init(struct kobject *kobj, struct kobj_type *ktype);
+
+The ktype is required for a kobject to be created properly, as every kobject
+must have an associated kobj_type.  After calling kobject_init(), to
+register the kobject with sysfs, the function kobject_add() must be called::
+
+    int kobject_add(struct kobject *kobj, struct kobject *parent,
+                    const char *fmt, ...);
+
+This sets up the parent of the kobject and the name for the kobject
+properly.  If the kobject is to be associated with a specific kset,
+kobj->kset must be assigned before calling kobject_add().  If a kset is
+associated with a kobject, then the parent for the kobject can be set to
+NULL in the call to kobject_add() and then the kobject's parent will be the
+kset itself.
+
+As the name of the kobject is set when it is added to the kernel, the name
+of the kobject should never be manipulated directly.  If you must change
+the name of the kobject, call kobject_rename()::
+
+    int kobject_rename(struct kobject *kobj, const char *new_name);
+
+kobject_rename() does not perform any locking or have a solid notion of
+what names are valid so the caller must provide their own sanity checking
+and serialization.
+
+There is a function called kobject_set_name() but that is legacy cruft and
+is being removed.  If your code needs to call this function, it is
+incorrect and needs to be fixed.
+
+To properly access the name of the kobject, use the function
+kobject_name()::
+
+    const char *kobject_name(const struct kobject * kobj);
+
+There is a helper function to both initialize and add the kobject to the
+kernel at the same time, called surprisingly enough kobject_init_and_add()::
+
+    int kobject_init_and_add(struct kobject *kobj, struct kobj_type *ktype,
+                             struct kobject *parent, const char *fmt, ...);
+
+The arguments are the same as the individual kobject_init() and
+kobject_add() functions described above.
+
+
+Uevents
+=======
+
+After a kobject has been registered with the kobject core, you need to
+announce to the world that it has been created.  This can be done with a
+call to kobject_uevent()::
+
+    int kobject_uevent(struct kobject *kobj, enum kobject_action action);
+
+Use the **KOBJ_ADD** action for when the kobject is first added to the kernel.
+This should be done only after any attributes or children of the kobject
+have been initialized properly, as userspace will instantly start to look
+for them when this call happens.
+
+When the kobject is removed from the kernel (details on how to do that are
+below), the uevent for **KOBJ_REMOVE** will be automatically created by the
+kobject core, so the caller does not have to worry about doing that by
+hand.
+
+
+Reference counts
+================
+
+One of the key functions of a kobject is to serve as a reference counter
+for the object in which it is embedded. As long as references to the object
+exist, the object (and the code which supports it) must continue to exist.
+The low-level functions for manipulating a kobject's reference counts are::
+
+    struct kobject *kobject_get(struct kobject *kobj);
+    void kobject_put(struct kobject *kobj);
+
+A successful call to kobject_get() will increment the kobject's reference
+counter and return the pointer to the kobject.
+
+When a reference is released, the call to kobject_put() will decrement the
+reference count and, possibly, free the object. Note that kobject_init()
+sets the reference count to one, so the code which sets up the kobject will
+need to do a kobject_put() eventually to release that reference.
+
+Because kobjects are dynamic, they must not be declared statically or on
+the stack, but instead, always allocated dynamically.  Future versions of
+the kernel will contain a run-time check for kobjects that are created
+statically and will warn the developer of this improper usage.
+
+If all that you want to use a kobject for is to provide a reference counter
+for your structure, please use the struct kref instead; a kobject would be
+overkill.  For more information on how to use struct kref, please see the
+file Documentation/core-api/kref.rst in the Linux kernel source tree.
+
+
+Creating "simple" kobjects
+==========================
+
+Sometimes all that a developer wants is a way to create a simple directory
+in the sysfs hierarchy, and not have to mess with the whole complication of
+ksets, show and store functions, and other details.  This is the one
+exception where a single kobject should be created.  To create such an
+entry, use the function::
+
+    struct kobject *kobject_create_and_add(const char *name, struct kobject *parent);
+
+This function will create a kobject and place it in sysfs in the location
+underneath the specified parent kobject.  To create simple attributes
+associated with this kobject, use::
+
+    int sysfs_create_file(struct kobject *kobj, const struct attribute *attr);
+
+or::
+
+    int sysfs_create_group(struct kobject *kobj, const struct attribute_group *grp);
+
+Both types of attributes used here, with a kobject that has been created
+with the kobject_create_and_add(), can be of type kobj_attribute, so no
+special custom attribute is needed to be created.
+
+See the example module, ``samples/kobject/kobject-example.c`` for an
+implementation of a simple kobject and attributes.
+
+
+
+ktypes and release methods
+==========================
+
+One important thing still missing from the discussion is what happens to a
+kobject when its reference count reaches zero. The code which created the
+kobject generally does not know when that will happen; if it did, there
+would be little point in using a kobject in the first place. Even
+predictable object lifecycles become more complicated when sysfs is brought
+in as other portions of the kernel can get a reference on any kobject that
+is registered in the system.
+
+The end result is that a structure protected by a kobject cannot be freed
+before its reference count goes to zero. The reference count is not under
+the direct control of the code which created the kobject. So that code must
+be notified asynchronously whenever the last reference to one of its
+kobjects goes away.
+
+Once you registered your kobject via kobject_add(), you must never use
+kfree() to free it directly. The only safe way is to use kobject_put(). It
+is good practice to always use kobject_put() after kobject_init() to avoid
+errors creeping in.
+
+This notification is done through a kobject's release() method. Usually
+such a method has a form like::
+
+    void my_object_release(struct kobject *kobj)
+    {
+            struct my_object *mine = container_of(kobj, struct my_object, kobj);
+
+            /* Perform any additional cleanup on this object, then... */
+            kfree(mine);
+    }
+
+One important point cannot be overstated: every kobject must have a
+release() method, and the kobject must persist (in a consistent state)
+until that method is called. If these constraints are not met, the code is
+flawed. Note that the kernel will warn you if you forget to provide a
+release() method.  Do not try to get rid of this warning by providing an
+"empty" release function.
+
+If all your cleanup function needs to do is call kfree(), then you must
+create a wrapper function which uses container_of() to upcast to the correct
+type (as shown in the example above) and then calls kfree() on the overall
+structure.
+
+Note, the name of the kobject is available in the release function, but it
+must NOT be changed within this callback.  Otherwise there will be a memory
+leak in the kobject core, which makes people unhappy.
+
+Interestingly, the release() method is not stored in the kobject itself;
+instead, it is associated with the ktype. So let us introduce struct
+kobj_type::
+
+    struct kobj_type {
+            void (*release)(struct kobject *kobj);
+            const struct sysfs_ops *sysfs_ops;
+            struct attribute **default_attrs;
+            const struct attribute_group **default_groups;
+            const struct kobj_ns_type_operations *(*child_ns_type)(struct kobject *kobj);
+            const void *(*namespace)(struct kobject *kobj);
+            void (*get_ownership)(struct kobject *kobj, kuid_t *uid, kgid_t *gid);
+    };
+
+This structure is used to describe a particular type of kobject (or, more
+correctly, of containing object). Every kobject needs to have an associated
+kobj_type structure; a pointer to that structure must be specified when you
+call kobject_init() or kobject_init_and_add().
+
+The release field in struct kobj_type is, of course, a pointer to the
+release() method for this type of kobject. The other two fields (sysfs_ops
+and default_attrs) control how objects of this type are represented in
+sysfs; they are beyond the scope of this document.
+
+The default_attrs pointer is a list of default attributes that will be
+automatically created for any kobject that is registered with this ktype.
+
+
+ksets
+=====
+
+A kset is merely a collection of kobjects that want to be associated with
+each other.  There is no restriction that they be of the same ktype, but be
+very careful if they are not.
+
+A kset serves these functions:
+
+ - It serves as a bag containing a group of objects. A kset can be used by
+   the kernel to track "all block devices" or "all PCI device drivers."
+
+ - A kset is also a subdirectory in sysfs, where the associated kobjects
+   with the kset can show up.  Every kset contains a kobject which can be
+   set up to be the parent of other kobjects; the top-level directories of
+   the sysfs hierarchy are constructed in this way.
+
+ - Ksets can support the "hotplugging" of kobjects and influence how
+   uevent events are reported to user space.
+
+In object-oriented terms, "kset" is the top-level container class; ksets
+contain their own kobject, but that kobject is managed by the kset code and
+should not be manipulated by any other user.
+
+A kset keeps its children in a standard kernel linked list.  Kobjects point
+back to their containing kset via their kset field. In almost all cases,
+the kobjects belonging to a kset have that kset (or, strictly, its embedded
+kobject) in their parent.
+
+As a kset contains a kobject within it, it should always be dynamically
+created and never declared statically or on the stack.  To create a new
+kset use::
+
+  struct kset *kset_create_and_add(const char *name,
+                                   const struct kset_uevent_ops *uevent_ops,
+                                   struct kobject *parent_kobj);
+
+When you are finished with the kset, call::
+
+  void kset_unregister(struct kset *k);
+
+to destroy it.  This removes the kset from sysfs and decrements its reference
+count.  When the reference count goes to zero, the kset will be released.
+Because other references to the kset may still exist, the release may happen
+after kset_unregister() returns.
+
+An example of using a kset can be seen in the
+``samples/kobject/kset-example.c`` file in the kernel tree.
+
+If a kset wishes to control the uevent operations of the kobjects
+associated with it, it can use the struct kset_uevent_ops to handle it::
+
+  struct kset_uevent_ops {
+          int (* const filter)(struct kset *kset, struct kobject *kobj);
+          const char *(* const name)(struct kset *kset, struct kobject *kobj);
+          int (* const uevent)(struct kset *kset, struct kobject *kobj,
+                        struct kobj_uevent_env *env);
+  };
+
+
+The filter function allows a kset to prevent a uevent from being emitted to
+userspace for a specific kobject.  If the function returns 0, the uevent
+will not be emitted.
+
+The name function will be called to override the default name of the kset
+that the uevent sends to userspace.  By default, the name will be the same
+as the kset itself, but this function, if present, can override that name.
+
+The uevent function will be called when the uevent is about to be sent to
+userspace to allow more environment variables to be added to the uevent.
+
+One might ask how, exactly, a kobject is added to a kset, given that no
+functions which perform that function have been presented.  The answer is
+that this task is handled by kobject_add().  When a kobject is passed to
+kobject_add(), its kset member should point to the kset to which the
+kobject will belong.  kobject_add() will handle the rest.
+
+If the kobject belonging to a kset has no parent kobject set, it will be
+added to the kset's directory.  Not all members of a kset do necessarily
+live in the kset directory.  If an explicit parent kobject is assigned
+before the kobject is added, the kobject is registered with the kset, but
+added below the parent kobject.
+
+
+Kobject removal
+===============
+
+After a kobject has been registered with the kobject core successfully, it
+must be cleaned up when the code is finished with it.  To do that, call
+kobject_put().  By doing this, the kobject core will automatically clean up
+all of the memory allocated by this kobject.  If a ``KOBJ_ADD`` uevent has been
+sent for the object, a corresponding ``KOBJ_REMOVE`` uevent will be sent, and
+any other sysfs housekeeping will be handled for the caller properly.
+
+If you need to do a two-stage delete of the kobject (say you are not
+allowed to sleep when you need to destroy the object), then call
+kobject_del() which will unregister the kobject from sysfs.  This makes the
+kobject "invisible", but it is not cleaned up, and the reference count of
+the object is still the same.  At a later time call kobject_put() to finish
+the cleanup of the memory associated with the kobject.
+
+kobject_del() can be used to drop the reference to the parent object, if
+circular references are constructed.  It is valid in some cases, that a
+parent objects references a child.  Circular references _must_ be broken
+with an explicit call to kobject_del(), so that a release functions will be
+called, and the objects in the former circle release each other.
+
+
+Example code to copy from
+=========================
+
+For a more complete example of using ksets and kobjects properly, see the
+example programs ``samples/kobject/{kobject-example.c,kset-example.c}``,
+which will be built as loadable modules if you select ``CONFIG_SAMPLE_KOBJECT``.
diff --git a/Documentation/core-api/kref.rst b/Documentation/core-api/kref.rst
new file mode 100644
index 000000000..c61eea6f1
--- /dev/null
+++ b/Documentation/core-api/kref.rst
@@ -0,0 +1,323 @@
+===================================================
+Adding reference counters (krefs) to kernel objects
+===================================================
+
+:Author: Corey Minyard <minyard@acm.org>
+:Author: Thomas Hellstrom <thellstrom@vmware.com>
+
+A lot of this was lifted from Greg Kroah-Hartman's 2004 OLS paper and
+presentation on krefs, which can be found at:
+
+  - http://www.kroah.com/linux/talks/ols_2004_kref_paper/Reprint-Kroah-Hartman-OLS2004.pdf
+  - http://www.kroah.com/linux/talks/ols_2004_kref_talk/
+
+Introduction
+============
+
+krefs allow you to add reference counters to your objects.  If you
+have objects that are used in multiple places and passed around, and
+you don't have refcounts, your code is almost certainly broken.  If
+you want refcounts, krefs are the way to go.
+
+To use a kref, add one to your data structures like::
+
+    struct my_data
+    {
+	.
+	.
+	struct kref refcount;
+	.
+	.
+    };
+
+The kref can occur anywhere within the data structure.
+
+Initialization
+==============
+
+You must initialize the kref after you allocate it.  To do this, call
+kref_init as so::
+
+     struct my_data *data;
+
+     data = kmalloc(sizeof(*data), GFP_KERNEL);
+     if (!data)
+            return -ENOMEM;
+     kref_init(&data->refcount);
+
+This sets the refcount in the kref to 1.
+
+Kref rules
+==========
+
+Once you have an initialized kref, you must follow the following
+rules:
+
+1) If you make a non-temporary copy of a pointer, especially if
+   it can be passed to another thread of execution, you must
+   increment the refcount with kref_get() before passing it off::
+
+       kref_get(&data->refcount);
+
+   If you already have a valid pointer to a kref-ed structure (the
+   refcount cannot go to zero) you may do this without a lock.
+
+2) When you are done with a pointer, you must call kref_put()::
+
+       kref_put(&data->refcount, data_release);
+
+   If this is the last reference to the pointer, the release
+   routine will be called.  If the code never tries to get
+   a valid pointer to a kref-ed structure without already
+   holding a valid pointer, it is safe to do this without
+   a lock.
+
+3) If the code attempts to gain a reference to a kref-ed structure
+   without already holding a valid pointer, it must serialize access
+   where a kref_put() cannot occur during the kref_get(), and the
+   structure must remain valid during the kref_get().
+
+For example, if you allocate some data and then pass it to another
+thread to process::
+
+    void data_release(struct kref *ref)
+    {
+	struct my_data *data = container_of(ref, struct my_data, refcount);
+	kfree(data);
+    }
+
+    void more_data_handling(void *cb_data)
+    {
+	struct my_data *data = cb_data;
+	.
+	. do stuff with data here
+	.
+	kref_put(&data->refcount, data_release);
+    }
+
+    int my_data_handler(void)
+    {
+	int rv = 0;
+	struct my_data *data;
+	struct task_struct *task;
+	data = kmalloc(sizeof(*data), GFP_KERNEL);
+	if (!data)
+		return -ENOMEM;
+	kref_init(&data->refcount);
+
+	kref_get(&data->refcount);
+	task = kthread_run(more_data_handling, data, "more_data_handling");
+	if (task == ERR_PTR(-ENOMEM)) {
+		rv = -ENOMEM;
+	        kref_put(&data->refcount, data_release);
+		goto out;
+	}
+
+	.
+	. do stuff with data here
+	.
+    out:
+	kref_put(&data->refcount, data_release);
+	return rv;
+    }
+
+This way, it doesn't matter what order the two threads handle the
+data, the kref_put() handles knowing when the data is not referenced
+any more and releasing it.  The kref_get() does not require a lock,
+since we already have a valid pointer that we own a refcount for.  The
+put needs no lock because nothing tries to get the data without
+already holding a pointer.
+
+In the above example, kref_put() will be called 2 times in both success
+and error paths. This is necessary because the reference count got
+incremented 2 times by kref_init() and kref_get().
+
+Note that the "before" in rule 1 is very important.  You should never
+do something like::
+
+	task = kthread_run(more_data_handling, data, "more_data_handling");
+	if (task == ERR_PTR(-ENOMEM)) {
+		rv = -ENOMEM;
+		goto out;
+	} else
+		/* BAD BAD BAD - get is after the handoff */
+		kref_get(&data->refcount);
+
+Don't assume you know what you are doing and use the above construct.
+First of all, you may not know what you are doing.  Second, you may
+know what you are doing (there are some situations where locking is
+involved where the above may be legal) but someone else who doesn't
+know what they are doing may change the code or copy the code.  It's
+bad style.  Don't do it.
+
+There are some situations where you can optimize the gets and puts.
+For instance, if you are done with an object and enqueuing it for
+something else or passing it off to something else, there is no reason
+to do a get then a put::
+
+	/* Silly extra get and put */
+	kref_get(&obj->ref);
+	enqueue(obj);
+	kref_put(&obj->ref, obj_cleanup);
+
+Just do the enqueue.  A comment about this is always welcome::
+
+	enqueue(obj);
+	/* We are done with obj, so we pass our refcount off
+	   to the queue.  DON'T TOUCH obj AFTER HERE! */
+
+The last rule (rule 3) is the nastiest one to handle.  Say, for
+instance, you have a list of items that are each kref-ed, and you wish
+to get the first one.  You can't just pull the first item off the list
+and kref_get() it.  That violates rule 3 because you are not already
+holding a valid pointer.  You must add a mutex (or some other lock).
+For instance::
+
+	static DEFINE_MUTEX(mutex);
+	static LIST_HEAD(q);
+	struct my_data
+	{
+		struct kref      refcount;
+		struct list_head link;
+	};
+
+	static struct my_data *get_entry()
+	{
+		struct my_data *entry = NULL;
+		mutex_lock(&mutex);
+		if (!list_empty(&q)) {
+			entry = container_of(q.next, struct my_data, link);
+			kref_get(&entry->refcount);
+		}
+		mutex_unlock(&mutex);
+		return entry;
+	}
+
+	static void release_entry(struct kref *ref)
+	{
+		struct my_data *entry = container_of(ref, struct my_data, refcount);
+
+		list_del(&entry->link);
+		kfree(entry);
+	}
+
+	static void put_entry(struct my_data *entry)
+	{
+		mutex_lock(&mutex);
+		kref_put(&entry->refcount, release_entry);
+		mutex_unlock(&mutex);
+	}
+
+The kref_put() return value is useful if you do not want to hold the
+lock during the whole release operation.  Say you didn't want to call
+kfree() with the lock held in the example above (since it is kind of
+pointless to do so).  You could use kref_put() as follows::
+
+	static void release_entry(struct kref *ref)
+	{
+		/* All work is done after the return from kref_put(). */
+	}
+
+	static void put_entry(struct my_data *entry)
+	{
+		mutex_lock(&mutex);
+		if (kref_put(&entry->refcount, release_entry)) {
+			list_del(&entry->link);
+			mutex_unlock(&mutex);
+			kfree(entry);
+		} else
+			mutex_unlock(&mutex);
+	}
+
+This is really more useful if you have to call other routines as part
+of the free operations that could take a long time or might claim the
+same lock.  Note that doing everything in the release routine is still
+preferred as it is a little neater.
+
+The above example could also be optimized using kref_get_unless_zero() in
+the following way::
+
+	static struct my_data *get_entry()
+	{
+		struct my_data *entry = NULL;
+		mutex_lock(&mutex);
+		if (!list_empty(&q)) {
+			entry = container_of(q.next, struct my_data, link);
+			if (!kref_get_unless_zero(&entry->refcount))
+				entry = NULL;
+		}
+		mutex_unlock(&mutex);
+		return entry;
+	}
+
+	static void release_entry(struct kref *ref)
+	{
+		struct my_data *entry = container_of(ref, struct my_data, refcount);
+
+		mutex_lock(&mutex);
+		list_del(&entry->link);
+		mutex_unlock(&mutex);
+		kfree(entry);
+	}
+
+	static void put_entry(struct my_data *entry)
+	{
+		kref_put(&entry->refcount, release_entry);
+	}
+
+Which is useful to remove the mutex lock around kref_put() in put_entry(), but
+it's important that kref_get_unless_zero is enclosed in the same critical
+section that finds the entry in the lookup table,
+otherwise kref_get_unless_zero may reference already freed memory.
+Note that it is illegal to use kref_get_unless_zero without checking its
+return value. If you are sure (by already having a valid pointer) that
+kref_get_unless_zero() will return true, then use kref_get() instead.
+
+Krefs and RCU
+=============
+
+The function kref_get_unless_zero also makes it possible to use rcu
+locking for lookups in the above example::
+
+	struct my_data
+	{
+		struct rcu_head rhead;
+		.
+		struct kref refcount;
+		.
+		.
+	};
+
+	static struct my_data *get_entry_rcu()
+	{
+		struct my_data *entry = NULL;
+		rcu_read_lock();
+		if (!list_empty(&q)) {
+			entry = container_of(q.next, struct my_data, link);
+			if (!kref_get_unless_zero(&entry->refcount))
+				entry = NULL;
+		}
+		rcu_read_unlock();
+		return entry;
+	}
+
+	static void release_entry_rcu(struct kref *ref)
+	{
+		struct my_data *entry = container_of(ref, struct my_data, refcount);
+
+		mutex_lock(&mutex);
+		list_del_rcu(&entry->link);
+		mutex_unlock(&mutex);
+		kfree_rcu(entry, rhead);
+	}
+
+	static void put_entry(struct my_data *entry)
+	{
+		kref_put(&entry->refcount, release_entry_rcu);
+	}
+
+But note that the struct kref member needs to remain in valid memory for a
+rcu grace period after release_entry_rcu was called. That can be accomplished
+by using kfree_rcu(entry, rhead) as done above, or by calling synchronize_rcu()
+before using kfree, but note that synchronize_rcu() may sleep for a
+substantial amount of time.
diff --git a/Documentation/core-api/librs.rst b/Documentation/core-api/librs.rst
new file mode 100644
index 000000000..6010f5bc5
--- /dev/null
+++ b/Documentation/core-api/librs.rst
@@ -0,0 +1,212 @@
+==========================================
+Reed-Solomon Library Programming Interface
+==========================================
+
+:Author: Thomas Gleixner
+
+Introduction
+============
+
+The generic Reed-Solomon Library provides encoding, decoding and error
+correction functions.
+
+Reed-Solomon codes are used in communication and storage applications to
+ensure data integrity.
+
+This documentation is provided for developers who want to utilize the
+functions provided by the library.
+
+Known Bugs And Assumptions
+==========================
+
+None.
+
+Usage
+=====
+
+This chapter provides examples of how to use the library.
+
+Initializing
+------------
+
+The init function init_rs returns a pointer to an rs decoder structure,
+which holds the necessary information for encoding, decoding and error
+correction with the given polynomial. It either uses an existing
+matching decoder or creates a new one. On creation all the lookup tables
+for fast en/decoding are created. The function may take a while, so make
+sure not to call it in critical code paths.
+
+::
+
+    /* the Reed Solomon control structure */
+    static struct rs_control *rs_decoder;
+
+    /* Symbolsize is 10 (bits)
+     * Primitive polynomial is x^10+x^3+1
+     * first consecutive root is 0
+     * primitive element to generate roots = 1
+     * generator polynomial degree (number of roots) = 6
+     */
+    rs_decoder = init_rs (10, 0x409, 0, 1, 6);
+
+
+Encoding
+--------
+
+The encoder calculates the Reed-Solomon code over the given data length
+and stores the result in the parity buffer. Note that the parity buffer
+must be initialized before calling the encoder.
+
+The expanded data can be inverted on the fly by providing a non-zero
+inversion mask. The expanded data is XOR'ed with the mask. This is used
+e.g. for FLASH ECC, where the all 0xFF is inverted to an all 0x00. The
+Reed-Solomon code for all 0x00 is all 0x00. The code is inverted before
+storing to FLASH so it is 0xFF too. This prevents that reading from an
+erased FLASH results in ECC errors.
+
+The databytes are expanded to the given symbol size on the fly. There is
+no support for encoding continuous bitstreams with a symbol size != 8 at
+the moment. If it is necessary it should be not a big deal to implement
+such functionality.
+
+::
+
+    /* Parity buffer. Size = number of roots */
+    uint16_t par[6];
+    /* Initialize the parity buffer */
+    memset(par, 0, sizeof(par));
+    /* Encode 512 byte in data8. Store parity in buffer par */
+    encode_rs8 (rs_decoder, data8, 512, par, 0);
+
+
+Decoding
+--------
+
+The decoder calculates the syndrome over the given data length and the
+received parity symbols and corrects errors in the data.
+
+If a syndrome is available from a hardware decoder then the syndrome
+calculation is skipped.
+
+The correction of the data buffer can be suppressed by providing a
+correction pattern buffer and an error location buffer to the decoder.
+The decoder stores the calculated error location and the correction
+bitmask in the given buffers. This is useful for hardware decoders which
+use a weird bit ordering scheme.
+
+The databytes are expanded to the given symbol size on the fly. There is
+no support for decoding continuous bitstreams with a symbolsize != 8 at
+the moment. If it is necessary it should be not a big deal to implement
+such functionality.
+
+Decoding with syndrome calculation, direct data correction
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+::
+
+    /* Parity buffer. Size = number of roots */
+    uint16_t par[6];
+    uint8_t  data[512];
+    int numerr;
+    /* Receive data */
+    .....
+    /* Receive parity */
+    .....
+    /* Decode 512 byte in data8.*/
+    numerr = decode_rs8 (rs_decoder, data8, par, 512, NULL, 0, NULL, 0, NULL);
+
+
+Decoding with syndrome given by hardware decoder, direct data correction
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+::
+
+    /* Parity buffer. Size = number of roots */
+    uint16_t par[6], syn[6];
+    uint8_t  data[512];
+    int numerr;
+    /* Receive data */
+    .....
+    /* Receive parity */
+    .....
+    /* Get syndrome from hardware decoder */
+    .....
+    /* Decode 512 byte in data8.*/
+    numerr = decode_rs8 (rs_decoder, data8, par, 512, syn, 0, NULL, 0, NULL);
+
+
+Decoding with syndrome given by hardware decoder, no direct data correction.
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+Note: It's not necessary to give data and received parity to the
+decoder.
+
+::
+
+    /* Parity buffer. Size = number of roots */
+    uint16_t par[6], syn[6], corr[8];
+    uint8_t  data[512];
+    int numerr, errpos[8];
+    /* Receive data */
+    .....
+    /* Receive parity */
+    .....
+    /* Get syndrome from hardware decoder */
+    .....
+    /* Decode 512 byte in data8.*/
+    numerr = decode_rs8 (rs_decoder, NULL, NULL, 512, syn, 0, errpos, 0, corr);
+    for (i = 0; i < numerr; i++) {
+        do_error_correction_in_your_buffer(errpos[i], corr[i]);
+    }
+
+
+Cleanup
+-------
+
+The function free_rs frees the allocated resources, if the caller is
+the last user of the decoder.
+
+::
+
+    /* Release resources */
+    free_rs(rs_decoder);
+
+
+Structures
+==========
+
+This chapter contains the autogenerated documentation of the structures
+which are used in the Reed-Solomon Library and are relevant for a
+developer.
+
+.. kernel-doc:: include/linux/rslib.h
+   :internal:
+
+Public Functions Provided
+=========================
+
+This chapter contains the autogenerated documentation of the
+Reed-Solomon functions which are exported.
+
+.. kernel-doc:: lib/reed_solomon/reed_solomon.c
+   :export:
+
+Credits
+=======
+
+The library code for encoding and decoding was written by Phil Karn.
+
+::
+
+            Copyright 2002, Phil Karn, KA9Q
+            May be used under the terms of the GNU General Public License (GPL)
+
+
+The wrapper functions and interfaces are written by Thomas Gleixner.
+
+Many users have provided bugfixes, improvements and helping hands for
+testing. Thanks a lot.
+
+The following people have contributed to this document:
+
+Thomas Gleixner\ tglx@linutronix.de
diff --git a/Documentation/core-api/local_ops.rst b/Documentation/core-api/local_ops.rst
new file mode 100644
index 000000000..2ac3f9f29
--- /dev/null
+++ b/Documentation/core-api/local_ops.rst
@@ -0,0 +1,202 @@
+
+.. _local_ops:
+
+=================================================
+Semantics and Behavior of Local Atomic Operations
+=================================================
+
+:Author: Mathieu Desnoyers
+
+
+This document explains the purpose of the local atomic operations, how
+to implement them for any given architecture and shows how they can be used
+properly. It also stresses on the precautions that must be taken when reading
+those local variables across CPUs when the order of memory writes matters.
+
+.. note::
+
+    Note that ``local_t`` based operations are not recommended for general
+    kernel use. Please use the ``this_cpu`` operations instead unless there is
+    really a special purpose. Most uses of ``local_t`` in the kernel have been
+    replaced by ``this_cpu`` operations. ``this_cpu`` operations combine the
+    relocation with the ``local_t`` like semantics in a single instruction and
+    yield more compact and faster executing code.
+
+
+Purpose of local atomic operations
+==================================
+
+Local atomic operations are meant to provide fast and highly reentrant per CPU
+counters. They minimize the performance cost of standard atomic operations by
+removing the LOCK prefix and memory barriers normally required to synchronize
+across CPUs.
+
+Having fast per CPU atomic counters is interesting in many cases: it does not
+require disabling interrupts to protect from interrupt handlers and it permits
+coherent counters in NMI handlers. It is especially useful for tracing purposes
+and for various performance monitoring counters.
+
+Local atomic operations only guarantee variable modification atomicity wrt the
+CPU which owns the data. Therefore, care must taken to make sure that only one
+CPU writes to the ``local_t`` data. This is done by using per cpu data and
+making sure that we modify it from within a preemption safe context. It is
+however permitted to read ``local_t`` data from any CPU: it will then appear to
+be written out of order wrt other memory writes by the owner CPU.
+
+
+Implementation for a given architecture
+=======================================
+
+It can be done by slightly modifying the standard atomic operations: only
+their UP variant must be kept. It typically means removing LOCK prefix (on
+i386 and x86_64) and any SMP synchronization barrier. If the architecture does
+not have a different behavior between SMP and UP, including
+``asm-generic/local.h`` in your architecture's ``local.h`` is sufficient.
+
+The ``local_t`` type is defined as an opaque ``signed long`` by embedding an
+``atomic_long_t`` inside a structure. This is made so a cast from this type to
+a ``long`` fails. The definition looks like::
+
+    typedef struct { atomic_long_t a; } local_t;
+
+
+Rules to follow when using local atomic operations
+==================================================
+
+* Variables touched by local ops must be per cpu variables.
+* *Only* the CPU owner of these variables must write to them.
+* This CPU can use local ops from any context (process, irq, softirq, nmi, ...)
+  to update its ``local_t`` variables.
+* Preemption (or interrupts) must be disabled when using local ops in
+  process context to make sure the process won't be migrated to a
+  different CPU between getting the per-cpu variable and doing the
+  actual local op.
+* When using local ops in interrupt context, no special care must be
+  taken on a mainline kernel, since they will run on the local CPU with
+  preemption already disabled. I suggest, however, to explicitly
+  disable preemption anyway to make sure it will still work correctly on
+  -rt kernels.
+* Reading the local cpu variable will provide the current copy of the
+  variable.
+* Reads of these variables can be done from any CPU, because updates to
+  "``long``", aligned, variables are always atomic. Since no memory
+  synchronization is done by the writer CPU, an outdated copy of the
+  variable can be read when reading some *other* cpu's variables.
+
+
+How to use local atomic operations
+==================================
+
+::
+
+    #include <linux/percpu.h>
+    #include <asm/local.h>
+
+    static DEFINE_PER_CPU(local_t, counters) = LOCAL_INIT(0);
+
+
+Counting
+========
+
+Counting is done on all the bits of a signed long.
+
+In preemptible context, use ``get_cpu_var()`` and ``put_cpu_var()`` around
+local atomic operations: it makes sure that preemption is disabled around write
+access to the per cpu variable. For instance::
+
+    local_inc(&get_cpu_var(counters));
+    put_cpu_var(counters);
+
+If you are already in a preemption-safe context, you can use
+``this_cpu_ptr()`` instead::
+
+    local_inc(this_cpu_ptr(&counters));
+
+
+
+Reading the counters
+====================
+
+Those local counters can be read from foreign CPUs to sum the count. Note that
+the data seen by local_read across CPUs must be considered to be out of order
+relatively to other memory writes happening on the CPU that owns the data::
+
+    long sum = 0;
+    for_each_online_cpu(cpu)
+            sum += local_read(&per_cpu(counters, cpu));
+
+If you want to use a remote local_read to synchronize access to a resource
+between CPUs, explicit ``smp_wmb()`` and ``smp_rmb()`` memory barriers must be used
+respectively on the writer and the reader CPUs. It would be the case if you use
+the ``local_t`` variable as a counter of bytes written in a buffer: there should
+be a ``smp_wmb()`` between the buffer write and the counter increment and also a
+``smp_rmb()`` between the counter read and the buffer read.
+
+
+Here is a sample module which implements a basic per cpu counter using
+``local.h``::
+
+    /* test-local.c
+     *
+     * Sample module for local.h usage.
+     */
+
+
+    #include <asm/local.h>
+    #include <linux/module.h>
+    #include <linux/timer.h>
+
+    static DEFINE_PER_CPU(local_t, counters) = LOCAL_INIT(0);
+
+    static struct timer_list test_timer;
+
+    /* IPI called on each CPU. */
+    static void test_each(void *info)
+    {
+            /* Increment the counter from a non preemptible context */
+            printk("Increment on cpu %d\n", smp_processor_id());
+            local_inc(this_cpu_ptr(&counters));
+
+            /* This is what incrementing the variable would look like within a
+             * preemptible context (it disables preemption) :
+             *
+             * local_inc(&get_cpu_var(counters));
+             * put_cpu_var(counters);
+             */
+    }
+
+    static void do_test_timer(unsigned long data)
+    {
+            int cpu;
+
+            /* Increment the counters */
+            on_each_cpu(test_each, NULL, 1);
+            /* Read all the counters */
+            printk("Counters read from CPU %d\n", smp_processor_id());
+            for_each_online_cpu(cpu) {
+                    printk("Read : CPU %d, count %ld\n", cpu,
+                            local_read(&per_cpu(counters, cpu)));
+            }
+            mod_timer(&test_timer, jiffies + 1000);
+    }
+
+    static int __init test_init(void)
+    {
+            /* initialize the timer that will increment the counter */
+            timer_setup(&test_timer, do_test_timer, 0);
+            mod_timer(&test_timer, jiffies + 1);
+
+            return 0;
+    }
+
+    static void __exit test_exit(void)
+    {
+            del_timer_sync(&test_timer);
+    }
+
+    module_init(test_init);
+    module_exit(test_exit);
+
+    MODULE_LICENSE("GPL");
+    MODULE_AUTHOR("Mathieu Desnoyers");
+    MODULE_DESCRIPTION("Local Atomic Ops");
diff --git a/Documentation/core-api/memory-allocation.rst b/Documentation/core-api/memory-allocation.rst
new file mode 100644
index 000000000..4446a1ac3
--- /dev/null
+++ b/Documentation/core-api/memory-allocation.rst
@@ -0,0 +1,172 @@
+.. _memory_allocation:
+
+=======================
+Memory Allocation Guide
+=======================
+
+Linux provides a variety of APIs for memory allocation. You can
+allocate small chunks using `kmalloc` or `kmem_cache_alloc` families,
+large virtually contiguous areas using `vmalloc` and its derivatives,
+or you can directly request pages from the page allocator with
+`alloc_pages`. It is also possible to use more specialized allocators,
+for instance `cma_alloc` or `zs_malloc`.
+
+Most of the memory allocation APIs use GFP flags to express how that
+memory should be allocated. The GFP acronym stands for "get free
+pages", the underlying memory allocation function.
+
+Diversity of the allocation APIs combined with the numerous GFP flags
+makes the question "How should I allocate memory?" not that easy to
+answer, although very likely you should use
+
+::
+
+  kzalloc(<size>, GFP_KERNEL);
+
+Of course there are cases when other allocation APIs and different GFP
+flags must be used.
+
+Get Free Page flags
+===================
+
+The GFP flags control the allocators behavior. They tell what memory
+zones can be used, how hard the allocator should try to find free
+memory, whether the memory can be accessed by the userspace etc. The
+:ref:`Documentation/core-api/mm-api.rst <mm-api-gfp-flags>` provides
+reference documentation for the GFP flags and their combinations and
+here we briefly outline their recommended usage:
+
+  * Most of the time ``GFP_KERNEL`` is what you need. Memory for the
+    kernel data structures, DMAable memory, inode cache, all these and
+    many other allocations types can use ``GFP_KERNEL``. Note, that
+    using ``GFP_KERNEL`` implies ``GFP_RECLAIM``, which means that
+    direct reclaim may be triggered under memory pressure; the calling
+    context must be allowed to sleep.
+  * If the allocation is performed from an atomic context, e.g interrupt
+    handler, use ``GFP_NOWAIT``. This flag prevents direct reclaim and
+    IO or filesystem operations. Consequently, under memory pressure
+    ``GFP_NOWAIT`` allocation is likely to fail. Allocations which
+    have a reasonable fallback should be using ``GFP_NOWARN``.
+  * If you think that accessing memory reserves is justified and the kernel
+    will be stressed unless allocation succeeds, you may use ``GFP_ATOMIC``.
+  * Untrusted allocations triggered from userspace should be a subject
+    of kmem accounting and must have ``__GFP_ACCOUNT`` bit set. There
+    is the handy ``GFP_KERNEL_ACCOUNT`` shortcut for ``GFP_KERNEL``
+    allocations that should be accounted.
+  * Userspace allocations should use either of the ``GFP_USER``,
+    ``GFP_HIGHUSER`` or ``GFP_HIGHUSER_MOVABLE`` flags. The longer
+    the flag name the less restrictive it is.
+
+    ``GFP_HIGHUSER_MOVABLE`` does not require that allocated memory
+    will be directly accessible by the kernel and implies that the
+    data is movable.
+
+    ``GFP_HIGHUSER`` means that the allocated memory is not movable,
+    but it is not required to be directly accessible by the kernel. An
+    example may be a hardware allocation that maps data directly into
+    userspace but has no addressing limitations.
+
+    ``GFP_USER`` means that the allocated memory is not movable and it
+    must be directly accessible by the kernel.
+
+You may notice that quite a few allocations in the existing code
+specify ``GFP_NOIO`` or ``GFP_NOFS``. Historically, they were used to
+prevent recursion deadlocks caused by direct memory reclaim calling
+back into the FS or IO paths and blocking on already held
+resources. Since 4.12 the preferred way to address this issue is to
+use new scope APIs described in
+:ref:`Documentation/core-api/gfp_mask-from-fs-io.rst <gfp_mask_from_fs_io>`.
+
+Other legacy GFP flags are ``GFP_DMA`` and ``GFP_DMA32``. They are
+used to ensure that the allocated memory is accessible by hardware
+with limited addressing capabilities. So unless you are writing a
+driver for a device with such restrictions, avoid using these flags.
+And even with hardware with restrictions it is preferable to use
+`dma_alloc*` APIs.
+
+GFP flags and reclaim behavior
+------------------------------
+Memory allocations may trigger direct or background reclaim and it is
+useful to understand how hard the page allocator will try to satisfy that
+or another request.
+
+  * ``GFP_KERNEL & ~__GFP_RECLAIM`` - optimistic allocation without _any_
+    attempt to free memory at all. The most light weight mode which even
+    doesn't kick the background reclaim. Should be used carefully because it
+    might deplete the memory and the next user might hit the more aggressive
+    reclaim.
+
+  * ``GFP_KERNEL & ~__GFP_DIRECT_RECLAIM`` (or ``GFP_NOWAIT``)- optimistic
+    allocation without any attempt to free memory from the current
+    context but can wake kswapd to reclaim memory if the zone is below
+    the low watermark. Can be used from either atomic contexts or when
+    the request is a performance optimization and there is another
+    fallback for a slow path.
+
+  * ``(GFP_KERNEL|__GFP_HIGH) & ~__GFP_DIRECT_RECLAIM`` (aka ``GFP_ATOMIC``) -
+    non sleeping allocation with an expensive fallback so it can access
+    some portion of memory reserves. Usually used from interrupt/bottom-half
+    context with an expensive slow path fallback.
+
+  * ``GFP_KERNEL`` - both background and direct reclaim are allowed and the
+    **default** page allocator behavior is used. That means that not costly
+    allocation requests are basically no-fail but there is no guarantee of
+    that behavior so failures have to be checked properly by callers
+    (e.g. OOM killer victim is allowed to fail currently).
+
+  * ``GFP_KERNEL | __GFP_NORETRY`` - overrides the default allocator behavior
+    and all allocation requests fail early rather than cause disruptive
+    reclaim (one round of reclaim in this implementation). The OOM killer
+    is not invoked.
+
+  * ``GFP_KERNEL | __GFP_RETRY_MAYFAIL`` - overrides the default allocator
+    behavior and all allocation requests try really hard. The request
+    will fail if the reclaim cannot make any progress. The OOM killer
+    won't be triggered.
+
+  * ``GFP_KERNEL | __GFP_NOFAIL`` - overrides the default allocator behavior
+    and all allocation requests will loop endlessly until they succeed.
+    This might be really dangerous especially for larger orders.
+
+Selecting memory allocator
+==========================
+
+The most straightforward way to allocate memory is to use a function
+from the kmalloc() family. And, to be on the safe side it's best to use
+routines that set memory to zero, like kzalloc(). If you need to
+allocate memory for an array, there are kmalloc_array() and kcalloc()
+helpers. The helpers struct_size(), array_size() and array3_size() can
+be used to safely calculate object sizes without overflowing.
+
+The maximal size of a chunk that can be allocated with `kmalloc` is
+limited. The actual limit depends on the hardware and the kernel
+configuration, but it is a good practice to use `kmalloc` for objects
+smaller than page size.
+
+The address of a chunk allocated with `kmalloc` is aligned to at least
+ARCH_KMALLOC_MINALIGN bytes.  For sizes which are a power of two, the
+alignment is also guaranteed to be at least the respective size.
+
+For large allocations you can use vmalloc() and vzalloc(), or directly
+request pages from the page allocator. The memory allocated by `vmalloc`
+and related functions is not physically contiguous.
+
+If you are not sure whether the allocation size is too large for
+`kmalloc`, it is possible to use kvmalloc() and its derivatives. It will
+try to allocate memory with `kmalloc` and if the allocation fails it
+will be retried with `vmalloc`. There are restrictions on which GFP
+flags can be used with `kvmalloc`; please see kvmalloc_node() reference
+documentation. Note that `kvmalloc` may return memory that is not
+physically contiguous.
+
+If you need to allocate many identical objects you can use the slab
+cache allocator. The cache should be set up with kmem_cache_create() or
+kmem_cache_create_usercopy() before it can be used. The second function
+should be used if a part of the cache might be copied to the userspace.
+After the cache is created kmem_cache_alloc() and its convenience
+wrappers can allocate memory from that cache.
+
+When the allocated memory is no longer needed it must be freed. You can
+use kvfree() for the memory allocated with `kmalloc`, `vmalloc` and
+`kvmalloc`. The slab caches should be freed with kmem_cache_free(). And
+don't forget to destroy the cache with kmem_cache_destroy().
diff --git a/Documentation/core-api/memory-hotplug.rst b/Documentation/core-api/memory-hotplug.rst
new file mode 100644
index 000000000..de7467e48
--- /dev/null
+++ b/Documentation/core-api/memory-hotplug.rst
@@ -0,0 +1,125 @@
+.. _memory_hotplug:
+
+==============
+Memory hotplug
+==============
+
+Memory hotplug event notifier
+=============================
+
+Hotplugging events are sent to a notification queue.
+
+There are six types of notification defined in ``include/linux/memory.h``:
+
+MEM_GOING_ONLINE
+  Generated before new memory becomes available in order to be able to
+  prepare subsystems to handle memory. The page allocator is still unable
+  to allocate from the new memory.
+
+MEM_CANCEL_ONLINE
+  Generated if MEM_GOING_ONLINE fails.
+
+MEM_ONLINE
+  Generated when memory has successfully brought online. The callback may
+  allocate pages from the new memory.
+
+MEM_GOING_OFFLINE
+  Generated to begin the process of offlining memory. Allocations are no
+  longer possible from the memory but some of the memory to be offlined
+  is still in use. The callback can be used to free memory known to a
+  subsystem from the indicated memory block.
+
+MEM_CANCEL_OFFLINE
+  Generated if MEM_GOING_OFFLINE fails. Memory is available again from
+  the memory block that we attempted to offline.
+
+MEM_OFFLINE
+  Generated after offlining memory is complete.
+
+A callback routine can be registered by calling::
+
+  hotplug_memory_notifier(callback_func, priority)
+
+Callback functions with higher values of priority are called before callback
+functions with lower values.
+
+A callback function must have the following prototype::
+
+  int callback_func(
+    struct notifier_block *self, unsigned long action, void *arg);
+
+The first argument of the callback function (self) is a pointer to the block
+of the notifier chain that points to the callback function itself.
+The second argument (action) is one of the event types described above.
+The third argument (arg) passes a pointer of struct memory_notify::
+
+	struct memory_notify {
+		unsigned long start_pfn;
+		unsigned long nr_pages;
+		int status_change_nid_normal;
+		int status_change_nid_high;
+		int status_change_nid;
+	}
+
+- start_pfn is start_pfn of online/offline memory.
+- nr_pages is # of pages of online/offline memory.
+- status_change_nid_normal is set node id when N_NORMAL_MEMORY of nodemask
+  is (will be) set/clear, if this is -1, then nodemask status is not changed.
+- status_change_nid_high is set node id when N_HIGH_MEMORY of nodemask
+  is (will be) set/clear, if this is -1, then nodemask status is not changed.
+- status_change_nid is set node id when N_MEMORY of nodemask is (will be)
+  set/clear. It means a new(memoryless) node gets new memory by online and a
+  node loses all memory. If this is -1, then nodemask status is not changed.
+
+  If status_changed_nid* >= 0, callback should create/discard structures for the
+  node if necessary.
+
+The callback routine shall return one of the values
+NOTIFY_DONE, NOTIFY_OK, NOTIFY_BAD, NOTIFY_STOP
+defined in ``include/linux/notifier.h``
+
+NOTIFY_DONE and NOTIFY_OK have no effect on the further processing.
+
+NOTIFY_BAD is used as response to the MEM_GOING_ONLINE, MEM_GOING_OFFLINE,
+MEM_ONLINE, or MEM_OFFLINE action to cancel hotplugging. It stops
+further processing of the notification queue.
+
+NOTIFY_STOP stops further processing of the notification queue.
+
+Locking Internals
+=================
+
+When adding/removing memory that uses memory block devices (i.e. ordinary RAM),
+the device_hotplug_lock should be held to:
+
+- synchronize against online/offline requests (e.g. via sysfs). This way, memory
+  block devices can only be accessed (.online/.state attributes) by user
+  space once memory has been fully added. And when removing memory, we
+  know nobody is in critical sections.
+- synchronize against CPU hotplug and similar (e.g. relevant for ACPI and PPC)
+
+Especially, there is a possible lock inversion that is avoided using
+device_hotplug_lock when adding memory and user space tries to online that
+memory faster than expected:
+
+- device_online() will first take the device_lock(), followed by
+  mem_hotplug_lock
+- add_memory_resource() will first take the mem_hotplug_lock, followed by
+  the device_lock() (while creating the devices, during bus_add_device()).
+
+As the device is visible to user space before taking the device_lock(), this
+can result in a lock inversion.
+
+onlining/offlining of memory should be done via device_online()/
+device_offline() - to make sure it is properly synchronized to actions
+via sysfs. Holding device_hotplug_lock is advised (to e.g. protect online_type)
+
+When adding/removing/onlining/offlining memory or adding/removing
+heterogeneous/device memory, we should always hold the mem_hotplug_lock in
+write mode to serialise memory hotplug (e.g. access to global/zone
+variables).
+
+In addition, mem_hotplug_lock (in contrast to device_hotplug_lock) in read
+mode allows for a quite efficient get_online_mems/put_online_mems
+implementation, so code accessing memory can protect from that memory
+vanishing.
diff --git a/Documentation/core-api/mm-api.rst b/Documentation/core-api/mm-api.rst
new file mode 100644
index 000000000..2adffb3f7
--- /dev/null
+++ b/Documentation/core-api/mm-api.rst
@@ -0,0 +1,97 @@
+======================
+Memory Management APIs
+======================
+
+User Space Memory Access
+========================
+
+.. kernel-doc:: arch/x86/include/asm/uaccess.h
+   :internal:
+
+.. kernel-doc:: arch/x86/lib/usercopy_32.c
+   :export:
+
+.. kernel-doc:: mm/gup.c
+   :functions: get_user_pages_fast
+
+.. _mm-api-gfp-flags:
+
+Memory Allocation Controls
+==========================
+
+Functions which need to allocate memory often use GFP flags to express
+how that memory should be allocated. The GFP acronym stands for "get
+free pages", the underlying memory allocation function. Not every GFP
+flag is allowed to every function which may allocate memory. Most
+users will want to use a plain ``GFP_KERNEL``.
+
+.. kernel-doc:: include/linux/gfp.h
+   :doc: Page mobility and placement hints
+
+.. kernel-doc:: include/linux/gfp.h
+   :doc: Watermark modifiers
+
+.. kernel-doc:: include/linux/gfp.h
+   :doc: Reclaim modifiers
+
+.. kernel-doc:: include/linux/gfp.h
+   :doc: Useful GFP flag combinations
+
+The Slab Cache
+==============
+
+.. kernel-doc:: include/linux/slab.h
+   :internal:
+
+.. kernel-doc:: mm/slab.c
+   :export:
+
+.. kernel-doc:: mm/slab_common.c
+   :export:
+
+.. kernel-doc:: mm/util.c
+   :functions: kfree_const kvmalloc_node kvfree
+
+Virtually Contiguous Mappings
+=============================
+
+.. kernel-doc:: mm/vmalloc.c
+   :export:
+
+File Mapping and Page Cache
+===========================
+
+.. kernel-doc:: mm/readahead.c
+   :export:
+
+.. kernel-doc:: mm/filemap.c
+   :export:
+
+.. kernel-doc:: mm/page-writeback.c
+   :export:
+
+.. kernel-doc:: mm/truncate.c
+   :export:
+
+.. kernel-doc:: include/linux/pagemap.h
+   :internal:
+
+Memory pools
+============
+
+.. kernel-doc:: mm/mempool.c
+   :export:
+
+DMA pools
+=========
+
+.. kernel-doc:: mm/dmapool.c
+   :export:
+
+More Memory Management Functions
+================================
+
+.. kernel-doc:: mm/memory.c
+   :export:
+
+.. kernel-doc:: mm/page_alloc.c
diff --git a/Documentation/core-api/packing.rst b/Documentation/core-api/packing.rst
new file mode 100644
index 000000000..d8c341fe3
--- /dev/null
+++ b/Documentation/core-api/packing.rst
@@ -0,0 +1,166 @@
+================================================
+Generic bitfield packing and unpacking functions
+================================================
+
+Problem statement
+-----------------
+
+When working with hardware, one has to choose between several approaches of
+interfacing with it.
+One can memory-map a pointer to a carefully crafted struct over the hardware
+device's memory region, and access its fields as struct members (potentially
+declared as bitfields). But writing code this way would make it less portable,
+due to potential endianness mismatches between the CPU and the hardware device.
+Additionally, one has to pay close attention when translating register
+definitions from the hardware documentation into bit field indices for the
+structs. Also, some hardware (typically networking equipment) tends to group
+its register fields in ways that violate any reasonable word boundaries
+(sometimes even 64 bit ones). This creates the inconvenience of having to
+define "high" and "low" portions of register fields within the struct.
+A more robust alternative to struct field definitions would be to extract the
+required fields by shifting the appropriate number of bits. But this would
+still not protect from endianness mismatches, except if all memory accesses
+were performed byte-by-byte. Also the code can easily get cluttered, and the
+high-level idea might get lost among the many bit shifts required.
+Many drivers take the bit-shifting approach and then attempt to reduce the
+clutter with tailored macros, but more often than not these macros take
+shortcuts that still prevent the code from being truly portable.
+
+The solution
+------------
+
+This API deals with 2 basic operations:
+
+  - Packing a CPU-usable number into a memory buffer (with hardware
+    constraints/quirks)
+  - Unpacking a memory buffer (which has hardware constraints/quirks)
+    into a CPU-usable number.
+
+The API offers an abstraction over said hardware constraints and quirks,
+over CPU endianness and therefore between possible mismatches between
+the two.
+
+The basic unit of these API functions is the u64. From the CPU's
+perspective, bit 63 always means bit offset 7 of byte 7, albeit only
+logically. The question is: where do we lay this bit out in memory?
+
+The following examples cover the memory layout of a packed u64 field.
+The byte offsets in the packed buffer are always implicitly 0, 1, ... 7.
+What the examples show is where the logical bytes and bits sit.
+
+1. Normally (no quirks), we would do it like this:
+
+::
+
+  63 62 61 60 59 58 57 56 55 54 53 52 51 50 49 48 47 46 45 44 43 42 41 40 39 38 37 36 35 34 33 32
+  7                       6                       5                        4
+  31 30 29 28 27 26 25 24 23 22 21 20 19 18 17 16 15 14 13 12 11 10  9  8  7  6  5  4  3  2  1  0
+  3                       2                       1                        0
+
+That is, the MSByte (7) of the CPU-usable u64 sits at memory offset 0, and the
+LSByte (0) of the u64 sits at memory offset 7.
+This corresponds to what most folks would regard to as "big endian", where
+bit i corresponds to the number 2^i. This is also referred to in the code
+comments as "logical" notation.
+
+
+2. If QUIRK_MSB_ON_THE_RIGHT is set, we do it like this:
+
+::
+
+  56 57 58 59 60 61 62 63 48 49 50 51 52 53 54 55 40 41 42 43 44 45 46 47 32 33 34 35 36 37 38 39
+  7                       6                        5                       4
+  24 25 26 27 28 29 30 31 16 17 18 19 20 21 22 23  8  9 10 11 12 13 14 15  0  1  2  3  4  5  6  7
+  3                       2                        1                       0
+
+That is, QUIRK_MSB_ON_THE_RIGHT does not affect byte positioning, but
+inverts bit offsets inside a byte.
+
+
+3. If QUIRK_LITTLE_ENDIAN is set, we do it like this:
+
+::
+
+  39 38 37 36 35 34 33 32 47 46 45 44 43 42 41 40 55 54 53 52 51 50 49 48 63 62 61 60 59 58 57 56
+  4                       5                       6                       7
+  7  6  5  4  3  2  1  0  15 14 13 12 11 10  9  8 23 22 21 20 19 18 17 16 31 30 29 28 27 26 25 24
+  0                       1                       2                       3
+
+Therefore, QUIRK_LITTLE_ENDIAN means that inside the memory region, every
+byte from each 4-byte word is placed at its mirrored position compared to
+the boundary of that word.
+
+4. If QUIRK_MSB_ON_THE_RIGHT and QUIRK_LITTLE_ENDIAN are both set, we do it
+   like this:
+
+::
+
+  32 33 34 35 36 37 38 39 40 41 42 43 44 45 46 47 48 49 50 51 52 53 54 55 56 57 58 59 60 61 62 63
+  4                       5                       6                       7
+  0  1  2  3  4  5  6  7  8   9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27 28 29 30 31
+  0                       1                       2                       3
+
+
+5. If just QUIRK_LSW32_IS_FIRST is set, we do it like this:
+
+::
+
+  31 30 29 28 27 26 25 24 23 22 21 20 19 18 17 16 15 14 13 12 11 10  9  8  7  6  5  4  3  2  1  0
+  3                       2                       1                        0
+  63 62 61 60 59 58 57 56 55 54 53 52 51 50 49 48 47 46 45 44 43 42 41 40 39 38 37 36 35 34 33 32
+  7                       6                       5                        4
+
+In this case the 8 byte memory region is interpreted as follows: first
+4 bytes correspond to the least significant 4-byte word, next 4 bytes to
+the more significant 4-byte word.
+
+
+6. If QUIRK_LSW32_IS_FIRST and QUIRK_MSB_ON_THE_RIGHT are set, we do it like
+   this:
+
+::
+
+  24 25 26 27 28 29 30 31 16 17 18 19 20 21 22 23  8  9 10 11 12 13 14 15  0  1  2  3  4  5  6  7
+  3                       2                        1                       0
+  56 57 58 59 60 61 62 63 48 49 50 51 52 53 54 55 40 41 42 43 44 45 46 47 32 33 34 35 36 37 38 39
+  7                       6                        5                       4
+
+
+7. If QUIRK_LSW32_IS_FIRST and QUIRK_LITTLE_ENDIAN are set, it looks like
+   this:
+
+::
+
+  7  6  5  4  3  2  1  0  15 14 13 12 11 10  9  8 23 22 21 20 19 18 17 16 31 30 29 28 27 26 25 24
+  0                       1                       2                       3
+  39 38 37 36 35 34 33 32 47 46 45 44 43 42 41 40 55 54 53 52 51 50 49 48 63 62 61 60 59 58 57 56
+  4                       5                       6                       7
+
+
+8. If QUIRK_LSW32_IS_FIRST, QUIRK_LITTLE_ENDIAN and QUIRK_MSB_ON_THE_RIGHT
+   are set, it looks like this:
+
+::
+
+  0  1  2  3  4  5  6  7  8   9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25 26 27 28 29 30 31
+  0                       1                       2                       3
+  32 33 34 35 36 37 38 39 40 41 42 43 44 45 46 47 48 49 50 51 52 53 54 55 56 57 58 59 60 61 62 63
+  4                       5                       6                       7
+
+
+We always think of our offsets as if there were no quirk, and we translate
+them afterwards, before accessing the memory region.
+
+Intended use
+------------
+
+Drivers that opt to use this API first need to identify which of the above 3
+quirk combinations (for a total of 8) match what the hardware documentation
+describes. Then they should wrap the packing() function, creating a new
+xxx_packing() that calls it using the proper QUIRK_* one-hot bits set.
+
+The packing() function returns an int-encoded error code, which protects the
+programmer against incorrect API use.  The errors are not expected to occur
+durring runtime, therefore it is reasonable for xxx_packing() to return void
+and simply swallow those errors. Optionally it can dump stack or print the
+error description.
diff --git a/Documentation/core-api/padata.rst b/Documentation/core-api/padata.rst
new file mode 100644
index 000000000..35175710b
--- /dev/null
+++ b/Documentation/core-api/padata.rst
@@ -0,0 +1,178 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=======================================
+The padata parallel execution mechanism
+=======================================
+
+:Date: May 2020
+
+Padata is a mechanism by which the kernel can farm jobs out to be done in
+parallel on multiple CPUs while optionally retaining their ordering.
+
+It was originally developed for IPsec, which needs to perform encryption and
+decryption on large numbers of packets without reordering those packets.  This
+is currently the sole consumer of padata's serialized job support.
+
+Padata also supports multithreaded jobs, splitting up the job evenly while load
+balancing and coordinating between threads.
+
+Running Serialized Jobs
+=======================
+
+Initializing
+------------
+
+The first step in using padata to run serialized jobs is to set up a
+padata_instance structure for overall control of how jobs are to be run::
+
+    #include <linux/padata.h>
+
+    struct padata_instance *padata_alloc(const char *name);
+
+'name' simply identifies the instance.
+
+Then, complete padata initialization by allocating a padata_shell::
+
+   struct padata_shell *padata_alloc_shell(struct padata_instance *pinst);
+
+A padata_shell is used to submit a job to padata and allows a series of such
+jobs to be serialized independently.  A padata_instance may have one or more
+padata_shells associated with it, each allowing a separate series of jobs.
+
+Modifying cpumasks
+------------------
+
+The CPUs used to run jobs can be changed in two ways, programatically with
+padata_set_cpumask() or via sysfs.  The former is defined::
+
+    int padata_set_cpumask(struct padata_instance *pinst, int cpumask_type,
+			   cpumask_var_t cpumask);
+
+Here cpumask_type is one of PADATA_CPU_PARALLEL or PADATA_CPU_SERIAL, where a
+parallel cpumask describes which processors will be used to execute jobs
+submitted to this instance in parallel and a serial cpumask defines which
+processors are allowed to be used as the serialization callback processor.
+cpumask specifies the new cpumask to use.
+
+There may be sysfs files for an instance's cpumasks.  For example, pcrypt's
+live in /sys/kernel/pcrypt/<instance-name>.  Within an instance's directory
+there are two files, parallel_cpumask and serial_cpumask, and either cpumask
+may be changed by echoing a bitmask into the file, for example::
+
+    echo f > /sys/kernel/pcrypt/pencrypt/parallel_cpumask
+
+Reading one of these files shows the user-supplied cpumask, which may be
+different from the 'usable' cpumask.
+
+Padata maintains two pairs of cpumasks internally, the user-supplied cpumasks
+and the 'usable' cpumasks.  (Each pair consists of a parallel and a serial
+cpumask.)  The user-supplied cpumasks default to all possible CPUs on instance
+allocation and may be changed as above.  The usable cpumasks are always a
+subset of the user-supplied cpumasks and contain only the online CPUs in the
+user-supplied masks; these are the cpumasks padata actually uses.  So it is
+legal to supply a cpumask to padata that contains offline CPUs.  Once an
+offline CPU in the user-supplied cpumask comes online, padata is going to use
+it.
+
+Changing the CPU masks are expensive operations, so it should not be done with
+great frequency.
+
+Running A Job
+-------------
+
+Actually submitting work to the padata instance requires the creation of a
+padata_priv structure, which represents one job::
+
+    struct padata_priv {
+        /* Other stuff here... */
+	void                    (*parallel)(struct padata_priv *padata);
+	void                    (*serial)(struct padata_priv *padata);
+    };
+
+This structure will almost certainly be embedded within some larger
+structure specific to the work to be done.  Most of its fields are private to
+padata, but the structure should be zeroed at initialisation time, and the
+parallel() and serial() functions should be provided.  Those functions will
+be called in the process of getting the work done as we will see
+momentarily.
+
+The submission of the job is done with::
+
+    int padata_do_parallel(struct padata_shell *ps,
+		           struct padata_priv *padata, int *cb_cpu);
+
+The ps and padata structures must be set up as described above; cb_cpu
+points to the preferred CPU to be used for the final callback when the job is
+done; it must be in the current instance's CPU mask (if not the cb_cpu pointer
+is updated to point to the CPU actually chosen).  The return value from
+padata_do_parallel() is zero on success, indicating that the job is in
+progress. -EBUSY means that somebody, somewhere else is messing with the
+instance's CPU mask, while -EINVAL is a complaint about cb_cpu not being in the
+serial cpumask, no online CPUs in the parallel or serial cpumasks, or a stopped
+instance.
+
+Each job submitted to padata_do_parallel() will, in turn, be passed to
+exactly one call to the above-mentioned parallel() function, on one CPU, so
+true parallelism is achieved by submitting multiple jobs.  parallel() runs with
+software interrupts disabled and thus cannot sleep.  The parallel()
+function gets the padata_priv structure pointer as its lone parameter;
+information about the actual work to be done is probably obtained by using
+container_of() to find the enclosing structure.
+
+Note that parallel() has no return value; the padata subsystem assumes that
+parallel() will take responsibility for the job from this point.  The job
+need not be completed during this call, but, if parallel() leaves work
+outstanding, it should be prepared to be called again with a new job before
+the previous one completes.
+
+Serializing Jobs
+----------------
+
+When a job does complete, parallel() (or whatever function actually finishes
+the work) should inform padata of the fact with a call to::
+
+    void padata_do_serial(struct padata_priv *padata);
+
+At some point in the future, padata_do_serial() will trigger a call to the
+serial() function in the padata_priv structure.  That call will happen on
+the CPU requested in the initial call to padata_do_parallel(); it, too, is
+run with local software interrupts disabled.
+Note that this call may be deferred for a while since the padata code takes
+pains to ensure that jobs are completed in the order in which they were
+submitted.
+
+Destroying
+----------
+
+Cleaning up a padata instance predictably involves calling the two free
+functions that correspond to the allocation in reverse::
+
+    void padata_free_shell(struct padata_shell *ps);
+    void padata_free(struct padata_instance *pinst);
+
+It is the user's responsibility to ensure all outstanding jobs are complete
+before any of the above are called.
+
+Running Multithreaded Jobs
+==========================
+
+A multithreaded job has a main thread and zero or more helper threads, with the
+main thread participating in the job and then waiting until all helpers have
+finished.  padata splits the job into units called chunks, where a chunk is a
+piece of the job that one thread completes in one call to the thread function.
+
+A user has to do three things to run a multithreaded job.  First, describe the
+job by defining a padata_mt_job structure, which is explained in the Interface
+section.  This includes a pointer to the thread function, which padata will
+call each time it assigns a job chunk to a thread.  Then, define the thread
+function, which accepts three arguments, ``start``, ``end``, and ``arg``, where
+the first two delimit the range that the thread operates on and the last is a
+pointer to the job's shared state, if any.  Prepare the shared state, which is
+typically allocated on the main thread's stack.  Last, call
+padata_do_multithreaded(), which will return once the job is finished.
+
+Interface
+=========
+
+.. kernel-doc:: include/linux/padata.h
+.. kernel-doc:: kernel/padata.c
diff --git a/Documentation/core-api/pin_user_pages.rst b/Documentation/core-api/pin_user_pages.rst
new file mode 100644
index 000000000..7ca8c7bac
--- /dev/null
+++ b/Documentation/core-api/pin_user_pages.rst
@@ -0,0 +1,279 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+====================================================
+pin_user_pages() and related calls
+====================================================
+
+.. contents:: :local:
+
+Overview
+========
+
+This document describes the following functions::
+
+ pin_user_pages()
+ pin_user_pages_fast()
+ pin_user_pages_remote()
+
+Basic description of FOLL_PIN
+=============================
+
+FOLL_PIN and FOLL_LONGTERM are flags that can be passed to the get_user_pages*()
+("gup") family of functions. FOLL_PIN has significant interactions and
+interdependencies with FOLL_LONGTERM, so both are covered here.
+
+FOLL_PIN is internal to gup, meaning that it should not appear at the gup call
+sites. This allows the associated wrapper functions  (pin_user_pages*() and
+others) to set the correct combination of these flags, and to check for problems
+as well.
+
+FOLL_LONGTERM, on the other hand, *is* allowed to be set at the gup call sites.
+This is in order to avoid creating a large number of wrapper functions to cover
+all combinations of get*(), pin*(), FOLL_LONGTERM, and more. Also, the
+pin_user_pages*() APIs are clearly distinct from the get_user_pages*() APIs, so
+that's a natural dividing line, and a good point to make separate wrapper calls.
+In other words, use pin_user_pages*() for DMA-pinned pages, and
+get_user_pages*() for other cases. There are five cases described later on in
+this document, to further clarify that concept.
+
+FOLL_PIN and FOLL_GET are mutually exclusive for a given gup call. However,
+multiple threads and call sites are free to pin the same struct pages, via both
+FOLL_PIN and FOLL_GET. It's just the call site that needs to choose one or the
+other, not the struct page(s).
+
+The FOLL_PIN implementation is nearly the same as FOLL_GET, except that FOLL_PIN
+uses a different reference counting technique.
+
+FOLL_PIN is a prerequisite to FOLL_LONGTERM. Another way of saying that is,
+FOLL_LONGTERM is a specific case, more restrictive case of FOLL_PIN.
+
+Which flags are set by each wrapper
+===================================
+
+For these pin_user_pages*() functions, FOLL_PIN is OR'd in with whatever gup
+flags the caller provides. The caller is required to pass in a non-null struct
+pages* array, and the function then pins pages by incrementing each by a special
+value: GUP_PIN_COUNTING_BIAS.
+
+For huge pages (and in fact, any compound page of more than 2 pages), the
+GUP_PIN_COUNTING_BIAS scheme is not used. Instead, an exact form of pin counting
+is achieved, by using the 3rd struct page in the compound page. A new struct
+page field, hpage_pinned_refcount, has been added in order to support this.
+
+This approach for compound pages avoids the counting upper limit problems that
+are discussed below. Those limitations would have been aggravated severely by
+huge pages, because each tail page adds a refcount to the head page. And in
+fact, testing revealed that, without a separate hpage_pinned_refcount field,
+page overflows were seen in some huge page stress tests.
+
+This also means that huge pages and compound pages (of order > 1) do not suffer
+from the false positives problem that is mentioned below.::
+
+ Function
+ --------
+ pin_user_pages          FOLL_PIN is always set internally by this function.
+ pin_user_pages_fast     FOLL_PIN is always set internally by this function.
+ pin_user_pages_remote   FOLL_PIN is always set internally by this function.
+
+For these get_user_pages*() functions, FOLL_GET might not even be specified.
+Behavior is a little more complex than above. If FOLL_GET was *not* specified,
+but the caller passed in a non-null struct pages* array, then the function
+sets FOLL_GET for you, and proceeds to pin pages by incrementing the refcount
+of each page by +1.::
+
+ Function
+ --------
+ get_user_pages           FOLL_GET is sometimes set internally by this function.
+ get_user_pages_fast      FOLL_GET is sometimes set internally by this function.
+ get_user_pages_remote    FOLL_GET is sometimes set internally by this function.
+
+Tracking dma-pinned pages
+=========================
+
+Some of the key design constraints, and solutions, for tracking dma-pinned
+pages:
+
+* An actual reference count, per struct page, is required. This is because
+  multiple processes may pin and unpin a page.
+
+* False positives (reporting that a page is dma-pinned, when in fact it is not)
+  are acceptable, but false negatives are not.
+
+* struct page may not be increased in size for this, and all fields are already
+  used.
+
+* Given the above, we can overload the page->_refcount field by using, sort of,
+  the upper bits in that field for a dma-pinned count. "Sort of", means that,
+  rather than dividing page->_refcount into bit fields, we simple add a medium-
+  large value (GUP_PIN_COUNTING_BIAS, initially chosen to be 1024: 10 bits) to
+  page->_refcount. This provides fuzzy behavior: if a page has get_page() called
+  on it 1024 times, then it will appear to have a single dma-pinned count.
+  And again, that's acceptable.
+
+This also leads to limitations: there are only 31-10==21 bits available for a
+counter that increments 10 bits at a time.
+
+* Callers must specifically request "dma-pinned tracking of pages". In other
+  words, just calling get_user_pages() will not suffice; a new set of functions,
+  pin_user_page() and related, must be used.
+
+FOLL_PIN, FOLL_GET, FOLL_LONGTERM: when to use which flags
+==========================================================
+
+Thanks to Jan Kara, Vlastimil Babka and several other -mm people, for describing
+these categories:
+
+CASE 1: Direct IO (DIO)
+-----------------------
+There are GUP references to pages that are serving
+as DIO buffers. These buffers are needed for a relatively short time (so they
+are not "long term"). No special synchronization with page_mkclean() or
+munmap() is provided. Therefore, flags to set at the call site are: ::
+
+    FOLL_PIN
+
+...but rather than setting FOLL_PIN directly, call sites should use one of
+the pin_user_pages*() routines that set FOLL_PIN.
+
+CASE 2: RDMA
+------------
+There are GUP references to pages that are serving as DMA
+buffers. These buffers are needed for a long time ("long term"). No special
+synchronization with page_mkclean() or munmap() is provided. Therefore, flags
+to set at the call site are: ::
+
+    FOLL_PIN | FOLL_LONGTERM
+
+NOTE: Some pages, such as DAX pages, cannot be pinned with longterm pins. That's
+because DAX pages do not have a separate page cache, and so "pinning" implies
+locking down file system blocks, which is not (yet) supported in that way.
+
+CASE 3: MMU notifier registration, with or without page faulting hardware
+-------------------------------------------------------------------------
+Device drivers can pin pages via get_user_pages*(), and register for mmu
+notifier callbacks for the memory range. Then, upon receiving a notifier
+"invalidate range" callback , stop the device from using the range, and unpin
+the pages. There may be other possible schemes, such as for example explicitly
+synchronizing against pending IO, that accomplish approximately the same thing.
+
+Or, if the hardware supports replayable page faults, then the device driver can
+avoid pinning entirely (this is ideal), as follows: register for mmu notifier
+callbacks as above, but instead of stopping the device and unpinning in the
+callback, simply remove the range from the device's page tables.
+
+Either way, as long as the driver unpins the pages upon mmu notifier callback,
+then there is proper synchronization with both filesystem and mm
+(page_mkclean(), munmap(), etc). Therefore, neither flag needs to be set.
+
+CASE 4: Pinning for struct page manipulation only
+-------------------------------------------------
+If only struct page data (as opposed to the actual memory contents that a page
+is tracking) is affected, then normal GUP calls are sufficient, and neither flag
+needs to be set.
+
+CASE 5: Pinning in order to write to the data within the page
+-------------------------------------------------------------
+Even though neither DMA nor Direct IO is involved, just a simple case of "pin,
+write to a page's data, unpin" can cause a problem. Case 5 may be considered a
+superset of Case 1, plus Case 2, plus anything that invokes that pattern. In
+other words, if the code is neither Case 1 nor Case 2, it may still require
+FOLL_PIN, for patterns like this:
+
+Correct (uses FOLL_PIN calls):
+    pin_user_pages()
+    write to the data within the pages
+    unpin_user_pages()
+
+INCORRECT (uses FOLL_GET calls):
+    get_user_pages()
+    write to the data within the pages
+    put_page()
+
+page_maybe_dma_pinned(): the whole point of pinning
+===================================================
+
+The whole point of marking pages as "DMA-pinned" or "gup-pinned" is to be able
+to query, "is this page DMA-pinned?" That allows code such as page_mkclean()
+(and file system writeback code in general) to make informed decisions about
+what to do when a page cannot be unmapped due to such pins.
+
+What to do in those cases is the subject of a years-long series of discussions
+and debates (see the References at the end of this document). It's a TODO item
+here: fill in the details once that's worked out. Meanwhile, it's safe to say
+that having this available: ::
+
+        static inline bool page_maybe_dma_pinned(struct page *page)
+
+...is a prerequisite to solving the long-running gup+DMA problem.
+
+Another way of thinking about FOLL_GET, FOLL_PIN, and FOLL_LONGTERM
+===================================================================
+
+Another way of thinking about these flags is as a progression of restrictions:
+FOLL_GET is for struct page manipulation, without affecting the data that the
+struct page refers to. FOLL_PIN is a *replacement* for FOLL_GET, and is for
+short term pins on pages whose data *will* get accessed. As such, FOLL_PIN is
+a "more severe" form of pinning. And finally, FOLL_LONGTERM is an even more
+restrictive case that has FOLL_PIN as a prerequisite: this is for pages that
+will be pinned longterm, and whose data will be accessed.
+
+Unit testing
+============
+This file::
+
+ tools/testing/selftests/vm/gup_benchmark.c
+
+has the following new calls to exercise the new pin*() wrapper functions:
+
+* PIN_FAST_BENCHMARK (./gup_benchmark -a)
+* PIN_BENCHMARK (./gup_benchmark -b)
+
+You can monitor how many total dma-pinned pages have been acquired and released
+since the system was booted, via two new /proc/vmstat entries: ::
+
+    /proc/vmstat/nr_foll_pin_acquired
+    /proc/vmstat/nr_foll_pin_released
+
+Under normal conditions, these two values will be equal unless there are any
+long-term [R]DMA pins in place, or during pin/unpin transitions.
+
+* nr_foll_pin_acquired: This is the number of logical pins that have been
+  acquired since the system was powered on. For huge pages, the head page is
+  pinned once for each page (head page and each tail page) within the huge page.
+  This follows the same sort of behavior that get_user_pages() uses for huge
+  pages: the head page is refcounted once for each tail or head page in the huge
+  page, when get_user_pages() is applied to a huge page.
+
+* nr_foll_pin_released: The number of logical pins that have been released since
+  the system was powered on. Note that pages are released (unpinned) on a
+  PAGE_SIZE granularity, even if the original pin was applied to a huge page.
+  Becaused of the pin count behavior described above in "nr_foll_pin_acquired",
+  the accounting balances out, so that after doing this::
+
+    pin_user_pages(huge_page);
+    for (each page in huge_page)
+        unpin_user_page(page);
+
+...the following is expected::
+
+    nr_foll_pin_released == nr_foll_pin_acquired
+
+(...unless it was already out of balance due to a long-term RDMA pin being in
+place.)
+
+Other diagnostics
+=================
+
+dump_page() has been enhanced slightly, to handle these new counting fields, and
+to better report on compound pages in general. Specifically, for compound pages
+with order > 1, the exact (hpage_pinned_refcount) pincount is reported.
+
+References
+==========
+
+* `Some slow progress on get_user_pages() (Apr 2, 2019) <https://lwn.net/Articles/784574/>`_
+* `DMA and get_user_pages() (LPC: Dec 12, 2018) <https://lwn.net/Articles/774411/>`_
+* `The trouble with get_user_pages() (Apr 30, 2018) <https://lwn.net/Articles/753027/>`_
+* `LWN kernel index: get_user_pages() <https://lwn.net/Kernel/Index/#Memory_management-get_user_pages>`_
+
+John Hubbard, October, 2019
diff --git a/Documentation/core-api/printk-basics.rst b/Documentation/core-api/printk-basics.rst
new file mode 100644
index 000000000..965e4281e
--- /dev/null
+++ b/Documentation/core-api/printk-basics.rst
@@ -0,0 +1,115 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===========================
+Message logging with printk
+===========================
+
+printk() is one of the most widely known functions in the Linux kernel. It's the
+standard tool we have for printing messages and usually the most basic way of
+tracing and debugging. If you're familiar with printf(3) you can tell printk()
+is based on it, although it has some functional differences:
+
+  - printk() messages can specify a log level.
+
+  - the format string, while largely compatible with C99, doesn't follow the
+    exact same specification. It has some extensions and a few limitations
+    (no ``%n`` or floating point conversion specifiers). See :ref:`How to get
+    printk format specifiers right <printk-specifiers>`.
+
+All printk() messages are printed to the kernel log buffer, which is a ring
+buffer exported to userspace through /dev/kmsg. The usual way to read it is
+using ``dmesg``.
+
+printk() is typically used like this::
+
+  printk(KERN_INFO "Message: %s\n", arg);
+
+where ``KERN_INFO`` is the log level (note that it's concatenated to the format
+string, the log level is not a separate argument). The available log levels are:
+
++----------------+--------+-----------------------------------------------+
+| Name           | String |  Alias function                               |
++================+========+===============================================+
+| KERN_EMERG     | "0"    | pr_emerg()                                    |
++----------------+--------+-----------------------------------------------+
+| KERN_ALERT     | "1"    | pr_alert()                                    |
++----------------+--------+-----------------------------------------------+
+| KERN_CRIT      | "2"    | pr_crit()                                     |
++----------------+--------+-----------------------------------------------+
+| KERN_ERR       | "3"    | pr_err()                                      |
++----------------+--------+-----------------------------------------------+
+| KERN_WARNING   | "4"    | pr_warn()                                     |
++----------------+--------+-----------------------------------------------+
+| KERN_NOTICE    | "5"    | pr_notice()                                   |
++----------------+--------+-----------------------------------------------+
+| KERN_INFO      | "6"    | pr_info()                                     |
++----------------+--------+-----------------------------------------------+
+| KERN_DEBUG     | "7"    | pr_debug() and pr_devel() if DEBUG is defined |
++----------------+--------+-----------------------------------------------+
+| KERN_DEFAULT   | ""     |                                               |
++----------------+--------+-----------------------------------------------+
+| KERN_CONT      | "c"    | pr_cont()                                     |
++----------------+--------+-----------------------------------------------+
+
+
+The log level specifies the importance of a message. The kernel decides whether
+to show the message immediately (printing it to the current console) depending
+on its log level and the current *console_loglevel* (a kernel variable). If the
+message priority is higher (lower log level value) than the *console_loglevel*
+the message will be printed to the console.
+
+If the log level is omitted, the message is printed with ``KERN_DEFAULT``
+level.
+
+You can check the current *console_loglevel* with::
+
+  $ cat /proc/sys/kernel/printk
+  4        4        1        7
+
+The result shows the *current*, *default*, *minimum* and *boot-time-default* log
+levels.
+
+To change the current console_loglevel simply write the desired level to
+``/proc/sys/kernel/printk``. For example, to print all messages to the console::
+
+  # echo 8 > /proc/sys/kernel/printk
+
+Another way, using ``dmesg``::
+
+  # dmesg -n 5
+
+sets the console_loglevel to print KERN_WARNING (4) or more severe messages to
+console. See ``dmesg(1)`` for more information.
+
+As an alternative to printk() you can use the ``pr_*()`` aliases for
+logging. This family of macros embed the log level in the macro names. For
+example::
+
+  pr_info("Info message no. %d\n", msg_num);
+
+prints a ``KERN_INFO`` message.
+
+Besides being more concise than the equivalent printk() calls, they can use a
+common definition for the format string through the pr_fmt() macro. For
+instance, defining this at the top of a source file (before any ``#include``
+directive)::
+
+  #define pr_fmt(fmt) "%s:%s: " fmt, KBUILD_MODNAME, __func__
+
+would prefix every pr_*() message in that file with the module and function name
+that originated the message.
+
+For debugging purposes there are also two conditionally-compiled macros:
+pr_debug() and pr_devel(), which are compiled-out unless ``DEBUG`` (or
+also ``CONFIG_DYNAMIC_DEBUG`` in the case of pr_debug()) is defined.
+
+
+Function reference
+==================
+
+.. kernel-doc:: kernel/printk/printk.c
+   :functions: printk
+
+.. kernel-doc:: include/linux/printk.h
+   :functions: pr_emerg pr_alert pr_crit pr_err pr_warn pr_notice pr_info
+      pr_fmt pr_debug pr_devel pr_cont
diff --git a/Documentation/core-api/printk-formats.rst b/Documentation/core-api/printk-formats.rst
new file mode 100644
index 000000000..6d26c5c6a
--- /dev/null
+++ b/Documentation/core-api/printk-formats.rst
@@ -0,0 +1,574 @@
+=========================================
+How to get printk format specifiers right
+=========================================
+
+.. _printk-specifiers:
+
+:Author: Randy Dunlap <rdunlap@infradead.org>
+:Author: Andrew Murray <amurray@mpc-data.co.uk>
+
+
+Integer types
+=============
+
+::
+
+	If variable is of Type,		use printk format specifier:
+	------------------------------------------------------------
+		char			%d or %x
+		unsigned char		%u or %x
+		short int		%d or %x
+		unsigned short int	%u or %x
+		int			%d or %x
+		unsigned int		%u or %x
+		long			%ld or %lx
+		unsigned long		%lu or %lx
+		long long		%lld or %llx
+		unsigned long long	%llu or %llx
+		size_t			%zu or %zx
+		ssize_t			%zd or %zx
+		s8			%d or %x
+		u8			%u or %x
+		s16			%d or %x
+		u16			%u or %x
+		s32			%d or %x
+		u32			%u or %x
+		s64			%lld or %llx
+		u64			%llu or %llx
+
+
+If <type> is dependent on a config option for its size (e.g., sector_t,
+blkcnt_t) or is architecture-dependent for its size (e.g., tcflag_t), use a
+format specifier of its largest possible type and explicitly cast to it.
+
+Example::
+
+	printk("test: sector number/total blocks: %llu/%llu\n",
+		(unsigned long long)sector, (unsigned long long)blockcount);
+
+Reminder: sizeof() returns type size_t.
+
+The kernel's printf does not support %n. Floating point formats (%e, %f,
+%g, %a) are also not recognized, for obvious reasons. Use of any
+unsupported specifier or length qualifier results in a WARN and early
+return from vsnprintf().
+
+Pointer types
+=============
+
+A raw pointer value may be printed with %p which will hash the address
+before printing. The kernel also supports extended specifiers for printing
+pointers of different types.
+
+Some of the extended specifiers print the data on the given address instead
+of printing the address itself. In this case, the following error messages
+might be printed instead of the unreachable information::
+
+	(null)	 data on plain NULL address
+	(efault) data on invalid address
+	(einval) invalid data on a valid address
+
+Plain Pointers
+--------------
+
+::
+
+	%p	abcdef12 or 00000000abcdef12
+
+Pointers printed without a specifier extension (i.e unadorned %p) are
+hashed to prevent leaking information about the kernel memory layout. This
+has the added benefit of providing a unique identifier. On 64-bit machines
+the first 32 bits are zeroed. The kernel will print ``(ptrval)`` until it
+gathers enough entropy. If you *really* want the address see %px below.
+
+Error Pointers
+--------------
+
+::
+
+	%pe	-ENOSPC
+
+For printing error pointers (i.e. a pointer for which IS_ERR() is true)
+as a symbolic error name. Error values for which no symbolic name is
+known are printed in decimal, while a non-ERR_PTR passed as the
+argument to %pe gets treated as ordinary %p.
+
+Symbols/Function Pointers
+-------------------------
+
+::
+
+	%pS	versatile_init+0x0/0x110
+	%ps	versatile_init
+	%pSR	versatile_init+0x9/0x110
+		(with __builtin_extract_return_addr() translation)
+	%pB	prev_fn_of_versatile_init+0x88/0x88
+
+
+The ``S`` and ``s`` specifiers are used for printing a pointer in symbolic
+format. They result in the symbol name with (S) or without (s)
+offsets. If KALLSYMS are disabled then the symbol address is printed instead.
+
+The ``B`` specifier results in the symbol name with offsets and should be
+used when printing stack backtraces. The specifier takes into
+consideration the effect of compiler optimisations which may occur
+when tail-calls are used and marked with the noreturn GCC attribute.
+
+Probed Pointers from BPF / tracing
+----------------------------------
+
+::
+
+	%pks	kernel string
+	%pus	user string
+
+The ``k`` and ``u`` specifiers are used for printing prior probed memory from
+either kernel memory (k) or user memory (u). The subsequent ``s`` specifier
+results in printing a string. For direct use in regular vsnprintf() the (k)
+and (u) annotation is ignored, however, when used out of BPF's bpf_trace_printk(),
+for example, it reads the memory it is pointing to without faulting.
+
+Kernel Pointers
+---------------
+
+::
+
+	%pK	01234567 or 0123456789abcdef
+
+For printing kernel pointers which should be hidden from unprivileged
+users. The behaviour of %pK depends on the kptr_restrict sysctl - see
+Documentation/admin-guide/sysctl/kernel.rst for more details.
+
+Unmodified Addresses
+--------------------
+
+::
+
+	%px	01234567 or 0123456789abcdef
+
+For printing pointers when you *really* want to print the address. Please
+consider whether or not you are leaking sensitive information about the
+kernel memory layout before printing pointers with %px. %px is functionally
+equivalent to %lx (or %lu). %px is preferred because it is more uniquely
+grep'able. If in the future we need to modify the way the kernel handles
+printing pointers we will be better equipped to find the call sites.
+
+Pointer Differences
+-------------------
+
+::
+
+	%td	2560
+	%tx	a00
+
+For printing the pointer differences, use the %t modifier for ptrdiff_t.
+
+Example::
+
+	printk("test: difference between pointers: %td\n", ptr2 - ptr1);
+
+Struct Resources
+----------------
+
+::
+
+	%pr	[mem 0x60000000-0x6fffffff flags 0x2200] or
+		[mem 0x0000000060000000-0x000000006fffffff flags 0x2200]
+	%pR	[mem 0x60000000-0x6fffffff pref] or
+		[mem 0x0000000060000000-0x000000006fffffff pref]
+
+For printing struct resources. The ``R`` and ``r`` specifiers result in a
+printed resource with (R) or without (r) a decoded flags member.
+
+Passed by reference.
+
+Physical address types phys_addr_t
+----------------------------------
+
+::
+
+	%pa[p]	0x01234567 or 0x0123456789abcdef
+
+For printing a phys_addr_t type (and its derivatives, such as
+resource_size_t) which can vary based on build options, regardless of the
+width of the CPU data path.
+
+Passed by reference.
+
+DMA address types dma_addr_t
+----------------------------
+
+::
+
+	%pad	0x01234567 or 0x0123456789abcdef
+
+For printing a dma_addr_t type which can vary based on build options,
+regardless of the width of the CPU data path.
+
+Passed by reference.
+
+Raw buffer as an escaped string
+-------------------------------
+
+::
+
+	%*pE[achnops]
+
+For printing raw buffer as an escaped string. For the following buffer::
+
+		1b 62 20 5c 43 07 22 90 0d 5d
+
+A few examples show how the conversion would be done (excluding surrounding
+quotes)::
+
+		%*pE		"\eb \C\a"\220\r]"
+		%*pEhp		"\x1bb \C\x07"\x90\x0d]"
+		%*pEa		"\e\142\040\\\103\a\042\220\r\135"
+
+The conversion rules are applied according to an optional combination
+of flags (see :c:func:`string_escape_mem` kernel documentation for the
+details):
+
+	- a - ESCAPE_ANY
+	- c - ESCAPE_SPECIAL
+	- h - ESCAPE_HEX
+	- n - ESCAPE_NULL
+	- o - ESCAPE_OCTAL
+	- p - ESCAPE_NP
+	- s - ESCAPE_SPACE
+
+By default ESCAPE_ANY_NP is used.
+
+ESCAPE_ANY_NP is the sane choice for many cases, in particularly for
+printing SSIDs.
+
+If field width is omitted then 1 byte only will be escaped.
+
+Raw buffer as a hex string
+--------------------------
+
+::
+
+	%*ph	00 01 02  ...  3f
+	%*phC	00:01:02: ... :3f
+	%*phD	00-01-02- ... -3f
+	%*phN	000102 ... 3f
+
+For printing small buffers (up to 64 bytes long) as a hex string with a
+certain separator. For larger buffers consider using
+:c:func:`print_hex_dump`.
+
+MAC/FDDI addresses
+------------------
+
+::
+
+	%pM	00:01:02:03:04:05
+	%pMR	05:04:03:02:01:00
+	%pMF	00-01-02-03-04-05
+	%pm	000102030405
+	%pmR	050403020100
+
+For printing 6-byte MAC/FDDI addresses in hex notation. The ``M`` and ``m``
+specifiers result in a printed address with (M) or without (m) byte
+separators. The default byte separator is the colon (:).
+
+Where FDDI addresses are concerned the ``F`` specifier can be used after
+the ``M`` specifier to use dash (-) separators instead of the default
+separator.
+
+For Bluetooth addresses the ``R`` specifier shall be used after the ``M``
+specifier to use reversed byte order suitable for visual interpretation
+of Bluetooth addresses which are in the little endian order.
+
+Passed by reference.
+
+IPv4 addresses
+--------------
+
+::
+
+	%pI4	1.2.3.4
+	%pi4	001.002.003.004
+	%p[Ii]4[hnbl]
+
+For printing IPv4 dot-separated decimal addresses. The ``I4`` and ``i4``
+specifiers result in a printed address with (i4) or without (I4) leading
+zeros.
+
+The additional ``h``, ``n``, ``b``, and ``l`` specifiers are used to specify
+host, network, big or little endian order addresses respectively. Where
+no specifier is provided the default network/big endian order is used.
+
+Passed by reference.
+
+IPv6 addresses
+--------------
+
+::
+
+	%pI6	0001:0002:0003:0004:0005:0006:0007:0008
+	%pi6	00010002000300040005000600070008
+	%pI6c	1:2:3:4:5:6:7:8
+
+For printing IPv6 network-order 16-bit hex addresses. The ``I6`` and ``i6``
+specifiers result in a printed address with (I6) or without (i6)
+colon-separators. Leading zeros are always used.
+
+The additional ``c`` specifier can be used with the ``I`` specifier to
+print a compressed IPv6 address as described by
+https://tools.ietf.org/html/rfc5952
+
+Passed by reference.
+
+IPv4/IPv6 addresses (generic, with port, flowinfo, scope)
+---------------------------------------------------------
+
+::
+
+	%pIS	1.2.3.4		or 0001:0002:0003:0004:0005:0006:0007:0008
+	%piS	001.002.003.004	or 00010002000300040005000600070008
+	%pISc	1.2.3.4		or 1:2:3:4:5:6:7:8
+	%pISpc	1.2.3.4:12345	or [1:2:3:4:5:6:7:8]:12345
+	%p[Ii]S[pfschnbl]
+
+For printing an IP address without the need to distinguish whether it's of
+type AF_INET or AF_INET6. A pointer to a valid struct sockaddr,
+specified through ``IS`` or ``iS``, can be passed to this format specifier.
+
+The additional ``p``, ``f``, and ``s`` specifiers are used to specify port
+(IPv4, IPv6), flowinfo (IPv6) and scope (IPv6). Ports have a ``:`` prefix,
+flowinfo a ``/`` and scope a ``%``, each followed by the actual value.
+
+In case of an IPv6 address the compressed IPv6 address as described by
+https://tools.ietf.org/html/rfc5952 is being used if the additional
+specifier ``c`` is given. The IPv6 address is surrounded by ``[``, ``]`` in
+case of additional specifiers ``p``, ``f`` or ``s`` as suggested by
+https://tools.ietf.org/html/draft-ietf-6man-text-addr-representation-07
+
+In case of IPv4 addresses, the additional ``h``, ``n``, ``b``, and ``l``
+specifiers can be used as well and are ignored in case of an IPv6
+address.
+
+Passed by reference.
+
+Further examples::
+
+	%pISfc		1.2.3.4		or [1:2:3:4:5:6:7:8]/123456789
+	%pISsc		1.2.3.4		or [1:2:3:4:5:6:7:8]%1234567890
+	%pISpfc		1.2.3.4:12345	or [1:2:3:4:5:6:7:8]:12345/123456789
+
+UUID/GUID addresses
+-------------------
+
+::
+
+	%pUb	00010203-0405-0607-0809-0a0b0c0d0e0f
+	%pUB	00010203-0405-0607-0809-0A0B0C0D0E0F
+	%pUl	03020100-0504-0706-0809-0a0b0c0e0e0f
+	%pUL	03020100-0504-0706-0809-0A0B0C0E0E0F
+
+For printing 16-byte UUID/GUIDs addresses. The additional ``l``, ``L``,
+``b`` and ``B`` specifiers are used to specify a little endian order in
+lower (l) or upper case (L) hex notation - and big endian order in lower (b)
+or upper case (B) hex notation.
+
+Where no additional specifiers are used the default big endian
+order with lower case hex notation will be printed.
+
+Passed by reference.
+
+dentry names
+------------
+
+::
+
+	%pd{,2,3,4}
+	%pD{,2,3,4}
+
+For printing dentry name; if we race with :c:func:`d_move`, the name might
+be a mix of old and new ones, but it won't oops.  %pd dentry is a safer
+equivalent of %s dentry->d_name.name we used to use, %pd<n> prints ``n``
+last components.  %pD does the same thing for struct file.
+
+Passed by reference.
+
+block_device names
+------------------
+
+::
+
+	%pg	sda, sda1 or loop0p1
+
+For printing name of block_device pointers.
+
+struct va_format
+----------------
+
+::
+
+	%pV
+
+For printing struct va_format structures. These contain a format string
+and va_list as follows::
+
+	struct va_format {
+		const char *fmt;
+		va_list *va;
+	};
+
+Implements a "recursive vsnprintf".
+
+Do not use this feature without some mechanism to verify the
+correctness of the format string and va_list arguments.
+
+Passed by reference.
+
+Device tree nodes
+-----------------
+
+::
+
+	%pOF[fnpPcCF]
+
+
+For printing device tree node structures. Default behaviour is
+equivalent to %pOFf.
+
+	- f - device node full_name
+	- n - device node name
+	- p - device node phandle
+	- P - device node path spec (name + @unit)
+	- F - device node flags
+	- c - major compatible string
+	- C - full compatible string
+
+The separator when using multiple arguments is ':'
+
+Examples::
+
+	%pOF	/foo/bar@0			- Node full name
+	%pOFf	/foo/bar@0			- Same as above
+	%pOFfp	/foo/bar@0:10			- Node full name + phandle
+	%pOFfcF	/foo/bar@0:foo,device:--P-	- Node full name +
+	                                          major compatible string +
+						  node flags
+							D - dynamic
+							d - detached
+							P - Populated
+							B - Populated bus
+
+Passed by reference.
+
+Fwnode handles
+--------------
+
+::
+
+	%pfw[fP]
+
+For printing information on fwnode handles. The default is to print the full
+node name, including the path. The modifiers are functionally equivalent to
+%pOF above.
+
+	- f - full name of the node, including the path
+	- P - the name of the node including an address (if there is one)
+
+Examples (ACPI)::
+
+	%pfwf	\_SB.PCI0.CIO2.port@1.endpoint@0	- Full node name
+	%pfwP	endpoint@0				- Node name
+
+Examples (OF)::
+
+	%pfwf	/ocp@68000000/i2c@48072000/camera@10/port/endpoint - Full name
+	%pfwP	endpoint				- Node name
+
+Time and date
+-------------
+
+::
+
+	%pt[RT]			YYYY-mm-ddTHH:MM:SS
+	%pt[RT]d		YYYY-mm-dd
+	%pt[RT]t		HH:MM:SS
+	%pt[RT][dt][r]
+
+For printing date and time as represented by::
+
+	R  struct rtc_time structure
+	T  time64_t type
+
+in human readable format.
+
+By default year will be incremented by 1900 and month by 1.
+Use %pt[RT]r (raw) to suppress this behaviour.
+
+Passed by reference.
+
+struct clk
+----------
+
+::
+
+	%pC	pll1
+	%pCn	pll1
+
+For printing struct clk structures. %pC and %pCn print the name of the clock
+(Common Clock Framework) or a unique 32-bit ID (legacy clock framework).
+
+Passed by reference.
+
+bitmap and its derivatives such as cpumask and nodemask
+-------------------------------------------------------
+
+::
+
+	%*pb	0779
+	%*pbl	0,3-6,8-10
+
+For printing bitmap and its derivatives such as cpumask and nodemask,
+%*pb outputs the bitmap with field width as the number of bits and %*pbl
+output the bitmap as range list with field width as the number of bits.
+
+Passed by reference.
+
+Flags bitfields such as page flags, gfp_flags
+---------------------------------------------
+
+::
+
+	%pGp	referenced|uptodate|lru|active|private
+	%pGg	GFP_USER|GFP_DMA32|GFP_NOWARN
+	%pGv	read|exec|mayread|maywrite|mayexec|denywrite
+
+For printing flags bitfields as a collection of symbolic constants that
+would construct the value. The type of flags is given by the third
+character. Currently supported are [p]age flags, [v]ma_flags (both
+expect ``unsigned long *``) and [g]fp_flags (expects ``gfp_t *``). The flag
+names and print order depends on the particular	type.
+
+Note that this format should not be used directly in the
+:c:func:`TP_printk()` part of a tracepoint. Instead, use the show_*_flags()
+functions from <trace/events/mmflags.h>.
+
+Passed by reference.
+
+Network device features
+-----------------------
+
+::
+
+	%pNF	0x000000000000c000
+
+For printing netdev_features_t.
+
+Passed by reference.
+
+Thanks
+======
+
+If you add other %p extensions, please extend <lib/test_printf.c> with
+one or more test cases, if at all feasible.
+
+Thank you for your cooperation and attention.
diff --git a/Documentation/core-api/protection-keys.rst b/Documentation/core-api/protection-keys.rst
new file mode 100644
index 000000000..ec575e72d
--- /dev/null
+++ b/Documentation/core-api/protection-keys.rst
@@ -0,0 +1,100 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+======================
+Memory Protection Keys
+======================
+
+Memory Protection Keys for Userspace (PKU aka PKEYs) is a feature
+which is found on Intel's Skylake (and later) "Scalable Processor"
+Server CPUs. It will be available in future non-server Intel parts
+and future AMD processors.
+
+For anyone wishing to test or use this feature, it is available in
+Amazon's EC2 C5 instances and is known to work there using an Ubuntu
+17.04 image.
+
+Memory Protection Keys provides a mechanism for enforcing page-based
+protections, but without requiring modification of the page tables
+when an application changes protection domains.  It works by
+dedicating 4 previously ignored bits in each page table entry to a
+"protection key", giving 16 possible keys.
+
+There is also a new user-accessible register (PKRU) with two separate
+bits (Access Disable and Write Disable) for each key.  Being a CPU
+register, PKRU is inherently thread-local, potentially giving each
+thread a different set of protections from every other thread.
+
+There are two new instructions (RDPKRU/WRPKRU) for reading and writing
+to the new register.  The feature is only available in 64-bit mode,
+even though there is theoretically space in the PAE PTEs.  These
+permissions are enforced on data access only and have no effect on
+instruction fetches.
+
+Syscalls
+========
+
+There are 3 system calls which directly interact with pkeys::
+
+	int pkey_alloc(unsigned long flags, unsigned long init_access_rights)
+	int pkey_free(int pkey);
+	int pkey_mprotect(unsigned long start, size_t len,
+			  unsigned long prot, int pkey);
+
+Before a pkey can be used, it must first be allocated with
+pkey_alloc().  An application calls the WRPKRU instruction
+directly in order to change access permissions to memory covered
+with a key.  In this example WRPKRU is wrapped by a C function
+called pkey_set().
+::
+
+	int real_prot = PROT_READ|PROT_WRITE;
+	pkey = pkey_alloc(0, PKEY_DISABLE_WRITE);
+	ptr = mmap(NULL, PAGE_SIZE, PROT_NONE, MAP_ANONYMOUS|MAP_PRIVATE, -1, 0);
+	ret = pkey_mprotect(ptr, PAGE_SIZE, real_prot, pkey);
+	... application runs here
+
+Now, if the application needs to update the data at 'ptr', it can
+gain access, do the update, then remove its write access::
+
+	pkey_set(pkey, 0); // clear PKEY_DISABLE_WRITE
+	*ptr = foo; // assign something
+	pkey_set(pkey, PKEY_DISABLE_WRITE); // set PKEY_DISABLE_WRITE again
+
+Now when it frees the memory, it will also free the pkey since it
+is no longer in use::
+
+	munmap(ptr, PAGE_SIZE);
+	pkey_free(pkey);
+
+.. note:: pkey_set() is a wrapper for the RDPKRU and WRPKRU instructions.
+          An example implementation can be found in
+          tools/testing/selftests/x86/protection_keys.c.
+
+Behavior
+========
+
+The kernel attempts to make protection keys consistent with the
+behavior of a plain mprotect().  For instance if you do this::
+
+	mprotect(ptr, size, PROT_NONE);
+	something(ptr);
+
+you can expect the same effects with protection keys when doing this::
+
+	pkey = pkey_alloc(0, PKEY_DISABLE_WRITE | PKEY_DISABLE_READ);
+	pkey_mprotect(ptr, size, PROT_READ|PROT_WRITE, pkey);
+	something(ptr);
+
+That should be true whether something() is a direct access to 'ptr'
+like::
+
+	*ptr = foo;
+
+or when the kernel does the access on the application's behalf like
+with a read()::
+
+	read(fd, ptr, 1);
+
+The kernel will send a SIGSEGV in both cases, but si_code will be set
+to SEGV_PKERR when violating protection keys versus SEGV_ACCERR when
+the plain mprotect() permissions are violated.
diff --git a/Documentation/core-api/rbtree.rst b/Documentation/core-api/rbtree.rst
new file mode 100644
index 000000000..6b88837fb
--- /dev/null
+++ b/Documentation/core-api/rbtree.rst
@@ -0,0 +1,429 @@
+=================================
+Red-black Trees (rbtree) in Linux
+=================================
+
+
+:Date: January 18, 2007
+:Author: Rob Landley <rob@landley.net>
+
+What are red-black trees, and what are they for?
+------------------------------------------------
+
+Red-black trees are a type of self-balancing binary search tree, used for
+storing sortable key/value data pairs.  This differs from radix trees (which
+are used to efficiently store sparse arrays and thus use long integer indexes
+to insert/access/delete nodes) and hash tables (which are not kept sorted to
+be easily traversed in order, and must be tuned for a specific size and
+hash function where rbtrees scale gracefully storing arbitrary keys).
+
+Red-black trees are similar to AVL trees, but provide faster real-time bounded
+worst case performance for insertion and deletion (at most two rotations and
+three rotations, respectively, to balance the tree), with slightly slower
+(but still O(log n)) lookup time.
+
+To quote Linux Weekly News:
+
+    There are a number of red-black trees in use in the kernel.
+    The deadline and CFQ I/O schedulers employ rbtrees to
+    track requests; the packet CD/DVD driver does the same.
+    The high-resolution timer code uses an rbtree to organize outstanding
+    timer requests.  The ext3 filesystem tracks directory entries in a
+    red-black tree.  Virtual memory areas (VMAs) are tracked with red-black
+    trees, as are epoll file descriptors, cryptographic keys, and network
+    packets in the "hierarchical token bucket" scheduler.
+
+This document covers use of the Linux rbtree implementation.  For more
+information on the nature and implementation of Red Black Trees,  see:
+
+  Linux Weekly News article on red-black trees
+    https://lwn.net/Articles/184495/
+
+  Wikipedia entry on red-black trees
+    https://en.wikipedia.org/wiki/Red-black_tree
+
+Linux implementation of red-black trees
+---------------------------------------
+
+Linux's rbtree implementation lives in the file "lib/rbtree.c".  To use it,
+"#include <linux/rbtree.h>".
+
+The Linux rbtree implementation is optimized for speed, and thus has one
+less layer of indirection (and better cache locality) than more traditional
+tree implementations.  Instead of using pointers to separate rb_node and data
+structures, each instance of struct rb_node is embedded in the data structure
+it organizes.  And instead of using a comparison callback function pointer,
+users are expected to write their own tree search and insert functions
+which call the provided rbtree functions.  Locking is also left up to the
+user of the rbtree code.
+
+Creating a new rbtree
+---------------------
+
+Data nodes in an rbtree tree are structures containing a struct rb_node member::
+
+  struct mytype {
+  	struct rb_node node;
+  	char *keystring;
+  };
+
+When dealing with a pointer to the embedded struct rb_node, the containing data
+structure may be accessed with the standard container_of() macro.  In addition,
+individual members may be accessed directly via rb_entry(node, type, member).
+
+At the root of each rbtree is an rb_root structure, which is initialized to be
+empty via:
+
+  struct rb_root mytree = RB_ROOT;
+
+Searching for a value in an rbtree
+----------------------------------
+
+Writing a search function for your tree is fairly straightforward: start at the
+root, compare each value, and follow the left or right branch as necessary.
+
+Example::
+
+  struct mytype *my_search(struct rb_root *root, char *string)
+  {
+  	struct rb_node *node = root->rb_node;
+
+  	while (node) {
+  		struct mytype *data = container_of(node, struct mytype, node);
+		int result;
+
+		result = strcmp(string, data->keystring);
+
+		if (result < 0)
+  			node = node->rb_left;
+		else if (result > 0)
+  			node = node->rb_right;
+		else
+  			return data;
+	}
+	return NULL;
+  }
+
+Inserting data into an rbtree
+-----------------------------
+
+Inserting data in the tree involves first searching for the place to insert the
+new node, then inserting the node and rebalancing ("recoloring") the tree.
+
+The search for insertion differs from the previous search by finding the
+location of the pointer on which to graft the new node.  The new node also
+needs a link to its parent node for rebalancing purposes.
+
+Example::
+
+  int my_insert(struct rb_root *root, struct mytype *data)
+  {
+  	struct rb_node **new = &(root->rb_node), *parent = NULL;
+
+  	/* Figure out where to put new node */
+  	while (*new) {
+  		struct mytype *this = container_of(*new, struct mytype, node);
+  		int result = strcmp(data->keystring, this->keystring);
+
+		parent = *new;
+  		if (result < 0)
+  			new = &((*new)->rb_left);
+  		else if (result > 0)
+  			new = &((*new)->rb_right);
+  		else
+  			return FALSE;
+  	}
+
+  	/* Add new node and rebalance tree. */
+  	rb_link_node(&data->node, parent, new);
+  	rb_insert_color(&data->node, root);
+
+	return TRUE;
+  }
+
+Removing or replacing existing data in an rbtree
+------------------------------------------------
+
+To remove an existing node from a tree, call::
+
+  void rb_erase(struct rb_node *victim, struct rb_root *tree);
+
+Example::
+
+  struct mytype *data = mysearch(&mytree, "walrus");
+
+  if (data) {
+  	rb_erase(&data->node, &mytree);
+  	myfree(data);
+  }
+
+To replace an existing node in a tree with a new one with the same key, call::
+
+  void rb_replace_node(struct rb_node *old, struct rb_node *new,
+  			struct rb_root *tree);
+
+Replacing a node this way does not re-sort the tree: If the new node doesn't
+have the same key as the old node, the rbtree will probably become corrupted.
+
+Iterating through the elements stored in an rbtree (in sort order)
+------------------------------------------------------------------
+
+Four functions are provided for iterating through an rbtree's contents in
+sorted order.  These work on arbitrary trees, and should not need to be
+modified or wrapped (except for locking purposes)::
+
+  struct rb_node *rb_first(struct rb_root *tree);
+  struct rb_node *rb_last(struct rb_root *tree);
+  struct rb_node *rb_next(struct rb_node *node);
+  struct rb_node *rb_prev(struct rb_node *node);
+
+To start iterating, call rb_first() or rb_last() with a pointer to the root
+of the tree, which will return a pointer to the node structure contained in
+the first or last element in the tree.  To continue, fetch the next or previous
+node by calling rb_next() or rb_prev() on the current node.  This will return
+NULL when there are no more nodes left.
+
+The iterator functions return a pointer to the embedded struct rb_node, from
+which the containing data structure may be accessed with the container_of()
+macro, and individual members may be accessed directly via
+rb_entry(node, type, member).
+
+Example::
+
+  struct rb_node *node;
+  for (node = rb_first(&mytree); node; node = rb_next(node))
+	printk("key=%s\n", rb_entry(node, struct mytype, node)->keystring);
+
+Cached rbtrees
+--------------
+
+Computing the leftmost (smallest) node is quite a common task for binary
+search trees, such as for traversals or users relying on a the particular
+order for their own logic. To this end, users can use 'struct rb_root_cached'
+to optimize O(logN) rb_first() calls to a simple pointer fetch avoiding
+potentially expensive tree iterations. This is done at negligible runtime
+overhead for maintanence; albeit larger memory footprint.
+
+Similar to the rb_root structure, cached rbtrees are initialized to be
+empty via::
+
+  struct rb_root_cached mytree = RB_ROOT_CACHED;
+
+Cached rbtree is simply a regular rb_root with an extra pointer to cache the
+leftmost node. This allows rb_root_cached to exist wherever rb_root does,
+which permits augmented trees to be supported as well as only a few extra
+interfaces::
+
+  struct rb_node *rb_first_cached(struct rb_root_cached *tree);
+  void rb_insert_color_cached(struct rb_node *, struct rb_root_cached *, bool);
+  void rb_erase_cached(struct rb_node *node, struct rb_root_cached *);
+
+Both insert and erase calls have their respective counterpart of augmented
+trees::
+
+  void rb_insert_augmented_cached(struct rb_node *node, struct rb_root_cached *,
+				  bool, struct rb_augment_callbacks *);
+  void rb_erase_augmented_cached(struct rb_node *, struct rb_root_cached *,
+				 struct rb_augment_callbacks *);
+
+
+Support for Augmented rbtrees
+-----------------------------
+
+Augmented rbtree is an rbtree with "some" additional data stored in
+each node, where the additional data for node N must be a function of
+the contents of all nodes in the subtree rooted at N. This data can
+be used to augment some new functionality to rbtree. Augmented rbtree
+is an optional feature built on top of basic rbtree infrastructure.
+An rbtree user who wants this feature will have to call the augmentation
+functions with the user provided augmentation callback when inserting
+and erasing nodes.
+
+C files implementing augmented rbtree manipulation must include
+<linux/rbtree_augmented.h> instead of <linux/rbtree.h>. Note that
+linux/rbtree_augmented.h exposes some rbtree implementations details
+you are not expected to rely on; please stick to the documented APIs
+there and do not include <linux/rbtree_augmented.h> from header files
+either so as to minimize chances of your users accidentally relying on
+such implementation details.
+
+On insertion, the user must update the augmented information on the path
+leading to the inserted node, then call rb_link_node() as usual and
+rb_augment_inserted() instead of the usual rb_insert_color() call.
+If rb_augment_inserted() rebalances the rbtree, it will callback into
+a user provided function to update the augmented information on the
+affected subtrees.
+
+When erasing a node, the user must call rb_erase_augmented() instead of
+rb_erase(). rb_erase_augmented() calls back into user provided functions
+to updated the augmented information on affected subtrees.
+
+In both cases, the callbacks are provided through struct rb_augment_callbacks.
+3 callbacks must be defined:
+
+- A propagation callback, which updates the augmented value for a given
+  node and its ancestors, up to a given stop point (or NULL to update
+  all the way to the root).
+
+- A copy callback, which copies the augmented value for a given subtree
+  to a newly assigned subtree root.
+
+- A tree rotation callback, which copies the augmented value for a given
+  subtree to a newly assigned subtree root AND recomputes the augmented
+  information for the former subtree root.
+
+The compiled code for rb_erase_augmented() may inline the propagation and
+copy callbacks, which results in a large function, so each augmented rbtree
+user should have a single rb_erase_augmented() call site in order to limit
+compiled code size.
+
+
+Sample usage
+^^^^^^^^^^^^
+
+Interval tree is an example of augmented rb tree. Reference -
+"Introduction to Algorithms" by Cormen, Leiserson, Rivest and Stein.
+More details about interval trees:
+
+Classical rbtree has a single key and it cannot be directly used to store
+interval ranges like [lo:hi] and do a quick lookup for any overlap with a new
+lo:hi or to find whether there is an exact match for a new lo:hi.
+
+However, rbtree can be augmented to store such interval ranges in a structured
+way making it possible to do efficient lookup and exact match.
+
+This "extra information" stored in each node is the maximum hi
+(max_hi) value among all the nodes that are its descendants. This
+information can be maintained at each node just be looking at the node
+and its immediate children. And this will be used in O(log n) lookup
+for lowest match (lowest start address among all possible matches)
+with something like::
+
+  struct interval_tree_node *
+  interval_tree_first_match(struct rb_root *root,
+			    unsigned long start, unsigned long last)
+  {
+	struct interval_tree_node *node;
+
+	if (!root->rb_node)
+		return NULL;
+	node = rb_entry(root->rb_node, struct interval_tree_node, rb);
+
+	while (true) {
+		if (node->rb.rb_left) {
+			struct interval_tree_node *left =
+				rb_entry(node->rb.rb_left,
+					 struct interval_tree_node, rb);
+			if (left->__subtree_last >= start) {
+				/*
+				 * Some nodes in left subtree satisfy Cond2.
+				 * Iterate to find the leftmost such node N.
+				 * If it also satisfies Cond1, that's the match
+				 * we are looking for. Otherwise, there is no
+				 * matching interval as nodes to the right of N
+				 * can't satisfy Cond1 either.
+				 */
+				node = left;
+				continue;
+			}
+		}
+		if (node->start <= last) {		/* Cond1 */
+			if (node->last >= start)	/* Cond2 */
+				return node;	/* node is leftmost match */
+			if (node->rb.rb_right) {
+				node = rb_entry(node->rb.rb_right,
+					struct interval_tree_node, rb);
+				if (node->__subtree_last >= start)
+					continue;
+			}
+		}
+		return NULL;	/* No match */
+	}
+  }
+
+Insertion/removal are defined using the following augmented callbacks::
+
+  static inline unsigned long
+  compute_subtree_last(struct interval_tree_node *node)
+  {
+	unsigned long max = node->last, subtree_last;
+	if (node->rb.rb_left) {
+		subtree_last = rb_entry(node->rb.rb_left,
+			struct interval_tree_node, rb)->__subtree_last;
+		if (max < subtree_last)
+			max = subtree_last;
+	}
+	if (node->rb.rb_right) {
+		subtree_last = rb_entry(node->rb.rb_right,
+			struct interval_tree_node, rb)->__subtree_last;
+		if (max < subtree_last)
+			max = subtree_last;
+	}
+	return max;
+  }
+
+  static void augment_propagate(struct rb_node *rb, struct rb_node *stop)
+  {
+	while (rb != stop) {
+		struct interval_tree_node *node =
+			rb_entry(rb, struct interval_tree_node, rb);
+		unsigned long subtree_last = compute_subtree_last(node);
+		if (node->__subtree_last == subtree_last)
+			break;
+		node->__subtree_last = subtree_last;
+		rb = rb_parent(&node->rb);
+	}
+  }
+
+  static void augment_copy(struct rb_node *rb_old, struct rb_node *rb_new)
+  {
+	struct interval_tree_node *old =
+		rb_entry(rb_old, struct interval_tree_node, rb);
+	struct interval_tree_node *new =
+		rb_entry(rb_new, struct interval_tree_node, rb);
+
+	new->__subtree_last = old->__subtree_last;
+  }
+
+  static void augment_rotate(struct rb_node *rb_old, struct rb_node *rb_new)
+  {
+	struct interval_tree_node *old =
+		rb_entry(rb_old, struct interval_tree_node, rb);
+	struct interval_tree_node *new =
+		rb_entry(rb_new, struct interval_tree_node, rb);
+
+	new->__subtree_last = old->__subtree_last;
+	old->__subtree_last = compute_subtree_last(old);
+  }
+
+  static const struct rb_augment_callbacks augment_callbacks = {
+	augment_propagate, augment_copy, augment_rotate
+  };
+
+  void interval_tree_insert(struct interval_tree_node *node,
+			    struct rb_root *root)
+  {
+	struct rb_node **link = &root->rb_node, *rb_parent = NULL;
+	unsigned long start = node->start, last = node->last;
+	struct interval_tree_node *parent;
+
+	while (*link) {
+		rb_parent = *link;
+		parent = rb_entry(rb_parent, struct interval_tree_node, rb);
+		if (parent->__subtree_last < last)
+			parent->__subtree_last = last;
+		if (start < parent->start)
+			link = &parent->rb.rb_left;
+		else
+			link = &parent->rb.rb_right;
+	}
+
+	node->__subtree_last = last;
+	rb_link_node(&node->rb, rb_parent, link);
+	rb_insert_augmented(&node->rb, root, &augment_callbacks);
+  }
+
+  void interval_tree_remove(struct interval_tree_node *node,
+			    struct rb_root *root)
+  {
+	rb_erase_augmented(&node->rb, root, &augment_callbacks);
+  }
diff --git a/Documentation/core-api/refcount-vs-atomic.rst b/Documentation/core-api/refcount-vs-atomic.rst
new file mode 100644
index 000000000..79a009ce1
--- /dev/null
+++ b/Documentation/core-api/refcount-vs-atomic.rst
@@ -0,0 +1,168 @@
+===================================
+refcount_t API compared to atomic_t
+===================================
+
+.. contents:: :local:
+
+Introduction
+============
+
+The goal of refcount_t API is to provide a minimal API for implementing
+an object's reference counters. While a generic architecture-independent
+implementation from lib/refcount.c uses atomic operations underneath,
+there are a number of differences between some of the ``refcount_*()`` and
+``atomic_*()`` functions with regards to the memory ordering guarantees.
+This document outlines the differences and provides respective examples
+in order to help maintainers validate their code against the change in
+these memory ordering guarantees.
+
+The terms used through this document try to follow the formal LKMM defined in
+tools/memory-model/Documentation/explanation.txt.
+
+memory-barriers.txt and atomic_t.txt provide more background to the
+memory ordering in general and for atomic operations specifically.
+
+Relevant types of memory ordering
+=================================
+
+.. note:: The following section only covers some of the memory
+   ordering types that are relevant for the atomics and reference
+   counters and used through this document. For a much broader picture
+   please consult memory-barriers.txt document.
+
+In the absence of any memory ordering guarantees (i.e. fully unordered)
+atomics & refcounters only provide atomicity and
+program order (po) relation (on the same CPU). It guarantees that
+each ``atomic_*()`` and ``refcount_*()`` operation is atomic and instructions
+are executed in program order on a single CPU.
+This is implemented using READ_ONCE()/WRITE_ONCE() and
+compare-and-swap primitives.
+
+A strong (full) memory ordering guarantees that all prior loads and
+stores (all po-earlier instructions) on the same CPU are completed
+before any po-later instruction is executed on the same CPU.
+It also guarantees that all po-earlier stores on the same CPU
+and all propagated stores from other CPUs must propagate to all
+other CPUs before any po-later instruction is executed on the original
+CPU (A-cumulative property). This is implemented using smp_mb().
+
+A RELEASE memory ordering guarantees that all prior loads and
+stores (all po-earlier instructions) on the same CPU are completed
+before the operation. It also guarantees that all po-earlier
+stores on the same CPU and all propagated stores from other CPUs
+must propagate to all other CPUs before the release operation
+(A-cumulative property). This is implemented using
+smp_store_release().
+
+An ACQUIRE memory ordering guarantees that all post loads and
+stores (all po-later instructions) on the same CPU are
+completed after the acquire operation. It also guarantees that all
+po-later stores on the same CPU must propagate to all other CPUs
+after the acquire operation executes. This is implemented using
+smp_acquire__after_ctrl_dep().
+
+A control dependency (on success) for refcounters guarantees that
+if a reference for an object was successfully obtained (reference
+counter increment or addition happened, function returned true),
+then further stores are ordered against this operation.
+Control dependency on stores are not implemented using any explicit
+barriers, but rely on CPU not to speculate on stores. This is only
+a single CPU relation and provides no guarantees for other CPUs.
+
+
+Comparison of functions
+=======================
+
+case 1) - non-"Read/Modify/Write" (RMW) ops
+-------------------------------------------
+
+Function changes:
+
+ * atomic_set() --> refcount_set()
+ * atomic_read() --> refcount_read()
+
+Memory ordering guarantee changes:
+
+ * none (both fully unordered)
+
+
+case 2) - increment-based ops that return no value
+--------------------------------------------------
+
+Function changes:
+
+ * atomic_inc() --> refcount_inc()
+ * atomic_add() --> refcount_add()
+
+Memory ordering guarantee changes:
+
+ * none (both fully unordered)
+
+case 3) - decrement-based RMW ops that return no value
+------------------------------------------------------
+
+Function changes:
+
+ * atomic_dec() --> refcount_dec()
+
+Memory ordering guarantee changes:
+
+ * fully unordered --> RELEASE ordering
+
+
+case 4) - increment-based RMW ops that return a value
+-----------------------------------------------------
+
+Function changes:
+
+ * atomic_inc_not_zero() --> refcount_inc_not_zero()
+ * no atomic counterpart --> refcount_add_not_zero()
+
+Memory ordering guarantees changes:
+
+ * fully ordered --> control dependency on success for stores
+
+.. note:: We really assume here that necessary ordering is provided as a
+   result of obtaining pointer to the object!
+
+
+case 5) - generic dec/sub decrement-based RMW ops that return a value
+---------------------------------------------------------------------
+
+Function changes:
+
+ * atomic_dec_and_test() --> refcount_dec_and_test()
+ * atomic_sub_and_test() --> refcount_sub_and_test()
+
+Memory ordering guarantees changes:
+
+ * fully ordered --> RELEASE ordering + ACQUIRE ordering on success
+
+
+case 6) other decrement-based RMW ops that return a value
+---------------------------------------------------------
+
+Function changes:
+
+ * no atomic counterpart --> refcount_dec_if_one()
+ * ``atomic_add_unless(&var, -1, 1)`` --> ``refcount_dec_not_one(&var)``
+
+Memory ordering guarantees changes:
+
+ * fully ordered --> RELEASE ordering + control dependency
+
+.. note:: atomic_add_unless() only provides full order on success.
+
+
+case 7) - lock-based RMW
+------------------------
+
+Function changes:
+
+ * atomic_dec_and_lock() --> refcount_dec_and_lock()
+ * atomic_dec_and_mutex_lock() --> refcount_dec_and_mutex_lock()
+
+Memory ordering guarantees changes:
+
+ * fully ordered --> RELEASE ordering + control dependency + hold
+   spin_lock() on success
diff --git a/Documentation/core-api/symbol-namespaces.rst b/Documentation/core-api/symbol-namespaces.rst
new file mode 100644
index 000000000..9b76337f6
--- /dev/null
+++ b/Documentation/core-api/symbol-namespaces.rst
@@ -0,0 +1,157 @@
+=================
+Symbol Namespaces
+=================
+
+The following document describes how to use Symbol Namespaces to structure the
+export surface of in-kernel symbols exported through the family of
+EXPORT_SYMBOL() macros.
+
+.. Table of Contents
+
+	=== 1 Introduction
+	=== 2 How to define Symbol Namespaces
+	   --- 2.1 Using the EXPORT_SYMBOL macros
+	   --- 2.2 Using the DEFAULT_SYMBOL_NAMESPACE define
+	=== 3 How to use Symbols exported in Namespaces
+	=== 4 Loading Modules that use namespaced Symbols
+	=== 5 Automatically creating MODULE_IMPORT_NS statements
+
+1. Introduction
+===============
+
+Symbol Namespaces have been introduced as a means to structure the export
+surface of the in-kernel API. It allows subsystem maintainers to partition
+their exported symbols into separate namespaces. That is useful for
+documentation purposes (think of the SUBSYSTEM_DEBUG namespace) as well as for
+limiting the availability of a set of symbols for use in other parts of the
+kernel. As of today, modules that make use of symbols exported into namespaces,
+are required to import the namespace. Otherwise the kernel will, depending on
+its configuration, reject loading the module or warn about a missing import.
+
+2. How to define Symbol Namespaces
+==================================
+
+Symbols can be exported into namespace using different methods. All of them are
+changing the way EXPORT_SYMBOL and friends are instrumented to create ksymtab
+entries.
+
+2.1 Using the EXPORT_SYMBOL macros
+==================================
+
+In addition to the macros EXPORT_SYMBOL() and EXPORT_SYMBOL_GPL(), that allow
+exporting of kernel symbols to the kernel symbol table, variants of these are
+available to export symbols into a certain namespace: EXPORT_SYMBOL_NS() and
+EXPORT_SYMBOL_NS_GPL(). They take one additional argument: the namespace.
+Please note that due to macro expansion that argument needs to be a
+preprocessor symbol. E.g. to export the symbol `usb_stor_suspend` into the
+namespace `USB_STORAGE`, use::
+
+	EXPORT_SYMBOL_NS(usb_stor_suspend, USB_STORAGE);
+
+The corresponding ksymtab entry struct `kernel_symbol` will have the member
+`namespace` set accordingly. A symbol that is exported without a namespace will
+refer to `NULL`. There is no default namespace if none is defined. `modpost`
+and kernel/module.c make use the namespace at build time or module load time,
+respectively.
+
+2.2 Using the DEFAULT_SYMBOL_NAMESPACE define
+=============================================
+
+Defining namespaces for all symbols of a subsystem can be very verbose and may
+become hard to maintain. Therefore a default define (DEFAULT_SYMBOL_NAMESPACE)
+is been provided, that, if set, will become the default for all EXPORT_SYMBOL()
+and EXPORT_SYMBOL_GPL() macro expansions that do not specify a namespace.
+
+There are multiple ways of specifying this define and it depends on the
+subsystem and the maintainer's preference, which one to use. The first option
+is to define the default namespace in the `Makefile` of the subsystem. E.g. to
+export all symbols defined in usb-common into the namespace USB_COMMON, add a
+line like this to drivers/usb/common/Makefile::
+
+	ccflags-y += -DDEFAULT_SYMBOL_NAMESPACE=USB_COMMON
+
+That will affect all EXPORT_SYMBOL() and EXPORT_SYMBOL_GPL() statements. A
+symbol exported with EXPORT_SYMBOL_NS() while this definition is present, will
+still be exported into the namespace that is passed as the namespace argument
+as this argument has preference over a default symbol namespace.
+
+A second option to define the default namespace is directly in the compilation
+unit as preprocessor statement. The above example would then read::
+
+	#undef  DEFAULT_SYMBOL_NAMESPACE
+	#define DEFAULT_SYMBOL_NAMESPACE USB_COMMON
+
+within the corresponding compilation unit before any EXPORT_SYMBOL macro is
+used.
+
+3. How to use Symbols exported in Namespaces
+============================================
+
+In order to use symbols that are exported into namespaces, kernel modules need
+to explicitly import these namespaces. Otherwise the kernel might reject to
+load the module. The module code is required to use the macro MODULE_IMPORT_NS
+for the namespaces it uses symbols from. E.g. a module using the
+usb_stor_suspend symbol from above, needs to import the namespace USB_STORAGE
+using a statement like::
+
+	MODULE_IMPORT_NS(USB_STORAGE);
+
+This will create a `modinfo` tag in the module for each imported namespace.
+This has the side effect, that the imported namespaces of a module can be
+inspected with modinfo::
+
+	$ modinfo drivers/usb/storage/ums-karma.ko
+	[...]
+	import_ns:      USB_STORAGE
+	[...]
+
+
+It is advisable to add the MODULE_IMPORT_NS() statement close to other module
+metadata definitions like MODULE_AUTHOR() or MODULE_LICENSE(). Refer to section
+5. for a way to create missing import statements automatically.
+
+4. Loading Modules that use namespaced Symbols
+==============================================
+
+At module loading time (e.g. `insmod`), the kernel will check each symbol
+referenced from the module for its availability and whether the namespace it
+might be exported to has been imported by the module. The default behaviour of
+the kernel is to reject loading modules that don't specify sufficient imports.
+An error will be logged and loading will be failed with EINVAL. In order to
+allow loading of modules that don't satisfy this precondition, a configuration
+option is available: Setting MODULE_ALLOW_MISSING_NAMESPACE_IMPORTS=y will
+enable loading regardless, but will emit a warning.
+
+5. Automatically creating MODULE_IMPORT_NS statements
+=====================================================
+
+Missing namespaces imports can easily be detected at build time. In fact,
+modpost will emit a warning if a module uses a symbol from a namespace
+without importing it.
+MODULE_IMPORT_NS() statements will usually be added at a definite location
+(along with other module meta data). To make the life of module authors (and
+subsystem maintainers) easier, a script and make target is available to fixup
+missing imports. Fixing missing imports can be done with::
+
+	$ make nsdeps
+
+A typical scenario for module authors would be::
+
+	- write code that depends on a symbol from a not imported namespace
+	- `make`
+	- notice the warning of modpost telling about a missing import
+	- run `make nsdeps` to add the import to the correct code location
+
+For subsystem maintainers introducing a namespace, the steps are very similar.
+Again, `make nsdeps` will eventually add the missing namespace imports for
+in-tree modules::
+
+	- move or add symbols to a namespace (e.g. with EXPORT_SYMBOL_NS())
+	- `make` (preferably with an allmodconfig to cover all in-kernel
+	  modules)
+	- notice the warning of modpost telling about a missing import
+	- run `make nsdeps` to add the import to the correct code location
+
+You can also run nsdeps for external module builds. A typical usage is::
+
+	$ make -C <path_to_kernel_src> M=$PWD nsdeps
diff --git a/Documentation/core-api/this_cpu_ops.rst b/Documentation/core-api/this_cpu_ops.rst
new file mode 100644
index 000000000..5cb8b883a
--- /dev/null
+++ b/Documentation/core-api/this_cpu_ops.rst
@@ -0,0 +1,339 @@
+===================
+this_cpu operations
+===================
+
+:Author: Christoph Lameter, August 4th, 2014
+:Author: Pranith Kumar, Aug 2nd, 2014
+
+this_cpu operations are a way of optimizing access to per cpu
+variables associated with the *currently* executing processor. This is
+done through the use of segment registers (or a dedicated register where
+the cpu permanently stored the beginning of the per cpu	area for a
+specific processor).
+
+this_cpu operations add a per cpu variable offset to the processor
+specific per cpu base and encode that operation in the instruction
+operating on the per cpu variable.
+
+This means that there are no atomicity issues between the calculation of
+the offset and the operation on the data. Therefore it is not
+necessary to disable preemption or interrupts to ensure that the
+processor is not changed between the calculation of the address and
+the operation on the data.
+
+Read-modify-write operations are of particular interest. Frequently
+processors have special lower latency instructions that can operate
+without the typical synchronization overhead, but still provide some
+sort of relaxed atomicity guarantees. The x86, for example, can execute
+RMW (Read Modify Write) instructions like inc/dec/cmpxchg without the
+lock prefix and the associated latency penalty.
+
+Access to the variable without the lock prefix is not synchronized but
+synchronization is not necessary since we are dealing with per cpu
+data specific to the currently executing processor. Only the current
+processor should be accessing that variable and therefore there are no
+concurrency issues with other processors in the system.
+
+Please note that accesses by remote processors to a per cpu area are
+exceptional situations and may impact performance and/or correctness
+(remote write operations) of local RMW operations via this_cpu_*.
+
+The main use of the this_cpu operations has been to optimize counter
+operations.
+
+The following this_cpu() operations with implied preemption protection
+are defined. These operations can be used without worrying about
+preemption and interrupts::
+
+	this_cpu_read(pcp)
+	this_cpu_write(pcp, val)
+	this_cpu_add(pcp, val)
+	this_cpu_and(pcp, val)
+	this_cpu_or(pcp, val)
+	this_cpu_add_return(pcp, val)
+	this_cpu_xchg(pcp, nval)
+	this_cpu_cmpxchg(pcp, oval, nval)
+	this_cpu_cmpxchg_double(pcp1, pcp2, oval1, oval2, nval1, nval2)
+	this_cpu_sub(pcp, val)
+	this_cpu_inc(pcp)
+	this_cpu_dec(pcp)
+	this_cpu_sub_return(pcp, val)
+	this_cpu_inc_return(pcp)
+	this_cpu_dec_return(pcp)
+
+
+Inner working of this_cpu operations
+------------------------------------
+
+On x86 the fs: or the gs: segment registers contain the base of the
+per cpu area. It is then possible to simply use the segment override
+to relocate a per cpu relative address to the proper per cpu area for
+the processor. So the relocation to the per cpu base is encoded in the
+instruction via a segment register prefix.
+
+For example::
+
+	DEFINE_PER_CPU(int, x);
+	int z;
+
+	z = this_cpu_read(x);
+
+results in a single instruction::
+
+	mov ax, gs:[x]
+
+instead of a sequence of calculation of the address and then a fetch
+from that address which occurs with the per cpu operations. Before
+this_cpu_ops such sequence also required preempt disable/enable to
+prevent the kernel from moving the thread to a different processor
+while the calculation is performed.
+
+Consider the following this_cpu operation::
+
+	this_cpu_inc(x)
+
+The above results in the following single instruction (no lock prefix!)::
+
+	inc gs:[x]
+
+instead of the following operations required if there is no segment
+register::
+
+	int *y;
+	int cpu;
+
+	cpu = get_cpu();
+	y = per_cpu_ptr(&x, cpu);
+	(*y)++;
+	put_cpu();
+
+Note that these operations can only be used on per cpu data that is
+reserved for a specific processor. Without disabling preemption in the
+surrounding code this_cpu_inc() will only guarantee that one of the
+per cpu counters is correctly incremented. However, there is no
+guarantee that the OS will not move the process directly before or
+after the this_cpu instruction is executed. In general this means that
+the value of the individual counters for each processor are
+meaningless. The sum of all the per cpu counters is the only value
+that is of interest.
+
+Per cpu variables are used for performance reasons. Bouncing cache
+lines can be avoided if multiple processors concurrently go through
+the same code paths.  Since each processor has its own per cpu
+variables no concurrent cache line updates take place. The price that
+has to be paid for this optimization is the need to add up the per cpu
+counters when the value of a counter is needed.
+
+
+Special operations
+------------------
+
+::
+
+	y = this_cpu_ptr(&x)
+
+Takes the offset of a per cpu variable (&x !) and returns the address
+of the per cpu variable that belongs to the currently executing
+processor.  this_cpu_ptr avoids multiple steps that the common
+get_cpu/put_cpu sequence requires. No processor number is
+available. Instead, the offset of the local per cpu area is simply
+added to the per cpu offset.
+
+Note that this operation is usually used in a code segment when
+preemption has been disabled. The pointer is then used to
+access local per cpu data in a critical section. When preemption
+is re-enabled this pointer is usually no longer useful since it may
+no longer point to per cpu data of the current processor.
+
+
+Per cpu variables and offsets
+-----------------------------
+
+Per cpu variables have *offsets* to the beginning of the per cpu
+area. They do not have addresses although they look like that in the
+code. Offsets cannot be directly dereferenced. The offset must be
+added to a base pointer of a per cpu area of a processor in order to
+form a valid address.
+
+Therefore the use of x or &x outside of the context of per cpu
+operations is invalid and will generally be treated like a NULL
+pointer dereference.
+
+::
+
+	DEFINE_PER_CPU(int, x);
+
+In the context of per cpu operations the above implies that x is a per
+cpu variable. Most this_cpu operations take a cpu variable.
+
+::
+
+	int __percpu *p = &x;
+
+&x and hence p is the *offset* of a per cpu variable. this_cpu_ptr()
+takes the offset of a per cpu variable which makes this look a bit
+strange.
+
+
+Operations on a field of a per cpu structure
+--------------------------------------------
+
+Let's say we have a percpu structure::
+
+	struct s {
+		int n,m;
+	};
+
+	DEFINE_PER_CPU(struct s, p);
+
+
+Operations on these fields are straightforward::
+
+	this_cpu_inc(p.m)
+
+	z = this_cpu_cmpxchg(p.m, 0, 1);
+
+
+If we have an offset to struct s::
+
+	struct s __percpu *ps = &p;
+
+	this_cpu_dec(ps->m);
+
+	z = this_cpu_inc_return(ps->n);
+
+
+The calculation of the pointer may require the use of this_cpu_ptr()
+if we do not make use of this_cpu ops later to manipulate fields::
+
+	struct s *pp;
+
+	pp = this_cpu_ptr(&p);
+
+	pp->m--;
+
+	z = pp->n++;
+
+
+Variants of this_cpu ops
+------------------------
+
+this_cpu ops are interrupt safe. Some architectures do not support
+these per cpu local operations. In that case the operation must be
+replaced by code that disables interrupts, then does the operations
+that are guaranteed to be atomic and then re-enable interrupts. Doing
+so is expensive. If there are other reasons why the scheduler cannot
+change the processor we are executing on then there is no reason to
+disable interrupts. For that purpose the following __this_cpu operations
+are provided.
+
+These operations have no guarantee against concurrent interrupts or
+preemption. If a per cpu variable is not used in an interrupt context
+and the scheduler cannot preempt, then they are safe. If any interrupts
+still occur while an operation is in progress and if the interrupt too
+modifies the variable, then RMW actions can not be guaranteed to be
+safe::
+
+	__this_cpu_read(pcp)
+	__this_cpu_write(pcp, val)
+	__this_cpu_add(pcp, val)
+	__this_cpu_and(pcp, val)
+	__this_cpu_or(pcp, val)
+	__this_cpu_add_return(pcp, val)
+	__this_cpu_xchg(pcp, nval)
+	__this_cpu_cmpxchg(pcp, oval, nval)
+	__this_cpu_cmpxchg_double(pcp1, pcp2, oval1, oval2, nval1, nval2)
+	__this_cpu_sub(pcp, val)
+	__this_cpu_inc(pcp)
+	__this_cpu_dec(pcp)
+	__this_cpu_sub_return(pcp, val)
+	__this_cpu_inc_return(pcp)
+	__this_cpu_dec_return(pcp)
+
+
+Will increment x and will not fall-back to code that disables
+interrupts on platforms that cannot accomplish atomicity through
+address relocation and a Read-Modify-Write operation in the same
+instruction.
+
+
+&this_cpu_ptr(pp)->n vs this_cpu_ptr(&pp->n)
+--------------------------------------------
+
+The first operation takes the offset and forms an address and then
+adds the offset of the n field. This may result in two add
+instructions emitted by the compiler.
+
+The second one first adds the two offsets and then does the
+relocation.  IMHO the second form looks cleaner and has an easier time
+with (). The second form also is consistent with the way
+this_cpu_read() and friends are used.
+
+
+Remote access to per cpu data
+------------------------------
+
+Per cpu data structures are designed to be used by one cpu exclusively.
+If you use the variables as intended, this_cpu_ops() are guaranteed to
+be "atomic" as no other CPU has access to these data structures.
+
+There are special cases where you might need to access per cpu data
+structures remotely. It is usually safe to do a remote read access
+and that is frequently done to summarize counters. Remote write access
+something which could be problematic because this_cpu ops do not
+have lock semantics. A remote write may interfere with a this_cpu
+RMW operation.
+
+Remote write accesses to percpu data structures are highly discouraged
+unless absolutely necessary. Please consider using an IPI to wake up
+the remote CPU and perform the update to its per cpu area.
+
+To access per-cpu data structure remotely, typically the per_cpu_ptr()
+function is used::
+
+
+	DEFINE_PER_CPU(struct data, datap);
+
+	struct data *p = per_cpu_ptr(&datap, cpu);
+
+This makes it explicit that we are getting ready to access a percpu
+area remotely.
+
+You can also do the following to convert the datap offset to an address::
+
+	struct data *p = this_cpu_ptr(&datap);
+
+but, passing of pointers calculated via this_cpu_ptr to other cpus is
+unusual and should be avoided.
+
+Remote access are typically only for reading the status of another cpus
+per cpu data. Write accesses can cause unique problems due to the
+relaxed synchronization requirements for this_cpu operations.
+
+One example that illustrates some concerns with write operations is
+the following scenario that occurs because two per cpu variables
+share a cache-line but the relaxed synchronization is applied to
+only one process updating the cache-line.
+
+Consider the following example::
+
+
+	struct test {
+		atomic_t a;
+		int b;
+	};
+
+	DEFINE_PER_CPU(struct test, onecacheline);
+
+There is some concern about what would happen if the field 'a' is updated
+remotely from one processor and the local processor would use this_cpu ops
+to update field b. Care should be taken that such simultaneous accesses to
+data within the same cache line are avoided. Also costly synchronization
+may be necessary. IPIs are generally recommended in such scenarios instead
+of a remote write to the per cpu area of another processor.
+
+Even in cases where the remote writes are rare, please bear in
+mind that a remote write will evict the cache line from the processor
+that most likely will access it. If the processor wakes up and finds a
+missing local cache line of a per cpu area, its performance and hence
+the wake up times will be affected.
diff --git a/Documentation/core-api/timekeeping.rst b/Documentation/core-api/timekeeping.rst
new file mode 100644
index 000000000..729e24864
--- /dev/null
+++ b/Documentation/core-api/timekeeping.rst
@@ -0,0 +1,189 @@
+ktime accessors
+===============
+
+Device drivers can read the current time using ktime_get() and the many
+related functions declared in linux/timekeeping.h. As a rule of thumb,
+using an accessor with a shorter name is preferred over one with a longer
+name if both are equally fit for a particular use case.
+
+Basic ktime_t based interfaces
+------------------------------
+
+The recommended simplest form returns an opaque ktime_t, with variants
+that return time for different clock references:
+
+
+.. c:function:: ktime_t ktime_get( void )
+
+	CLOCK_MONOTONIC
+
+	Useful for reliable timestamps and measuring short time intervals
+	accurately. Starts at system boot time but stops during suspend.
+
+.. c:function:: ktime_t ktime_get_boottime( void )
+
+	CLOCK_BOOTTIME
+
+	Like ktime_get(), but does not stop when suspended. This can be
+	used e.g. for key expiration times that need to be synchronized
+	with other machines across a suspend operation.
+
+.. c:function:: ktime_t ktime_get_real( void )
+
+	CLOCK_REALTIME
+
+	Returns the time in relative to the UNIX epoch starting in 1970
+	using the Coordinated Universal Time (UTC), same as gettimeofday()
+	user space. This is used for all timestamps that need to
+	persist across a reboot, like inode times, but should be avoided
+	for internal uses, since it can jump backwards due to a leap
+	second update, NTP adjustment settimeofday() operation from user
+	space.
+
+.. c:function:: ktime_t ktime_get_clocktai( void )
+
+	 CLOCK_TAI
+
+	Like ktime_get_real(), but uses the International Atomic Time (TAI)
+	reference instead of UTC to avoid jumping on leap second updates.
+	This is rarely useful in the kernel.
+
+.. c:function:: ktime_t ktime_get_raw( void )
+
+	CLOCK_MONOTONIC_RAW
+
+	Like ktime_get(), but runs at the same rate as the hardware
+	clocksource without (NTP) adjustments for clock drift. This is
+	also rarely needed in the kernel.
+
+nanosecond, timespec64, and second output
+-----------------------------------------
+
+For all of the above, there are variants that return the time in a
+different format depending on what is required by the user:
+
+.. c:function:: u64 ktime_get_ns( void )
+		u64 ktime_get_boottime_ns( void )
+		u64 ktime_get_real_ns( void )
+		u64 ktime_get_clocktai_ns( void )
+		u64 ktime_get_raw_ns( void )
+
+	Same as the plain ktime_get functions, but returning a u64 number
+	of nanoseconds in the respective time reference, which may be
+	more convenient for some callers.
+
+.. c:function:: void ktime_get_ts64( struct timespec64 * )
+		void ktime_get_boottime_ts64( struct timespec64 * )
+		void ktime_get_real_ts64( struct timespec64 * )
+		void ktime_get_clocktai_ts64( struct timespec64 * )
+		void ktime_get_raw_ts64( struct timespec64 * )
+
+	Same above, but returns the time in a 'struct timespec64', split
+	into seconds and nanoseconds. This can avoid an extra division
+	when printing the time, or when passing it into an external
+	interface that expects a 'timespec' or 'timeval' structure.
+
+.. c:function:: time64_t ktime_get_seconds( void )
+		time64_t ktime_get_boottime_seconds( void )
+		time64_t ktime_get_real_seconds( void )
+		time64_t ktime_get_clocktai_seconds( void )
+		time64_t ktime_get_raw_seconds( void )
+
+	Return a coarse-grained version of the time as a scalar
+	time64_t. This avoids accessing the clock hardware and rounds
+	down the seconds to the full seconds of the last timer tick
+	using the respective reference.
+
+Coarse and fast_ns access
+-------------------------
+
+Some additional variants exist for more specialized cases:
+
+.. c:function:: ktime_t ktime_get_coarse( void )
+		ktime_t ktime_get_coarse_boottime( void )
+		ktime_t ktime_get_coarse_real( void )
+		ktime_t ktime_get_coarse_clocktai( void )
+
+.. c:function:: u64 ktime_get_coarse_ns( void )
+		u64 ktime_get_coarse_boottime_ns( void )
+		u64 ktime_get_coarse_real_ns( void )
+		u64 ktime_get_coarse_clocktai_ns( void )
+
+.. c:function:: void ktime_get_coarse_ts64( struct timespec64 * )
+		void ktime_get_coarse_boottime_ts64( struct timespec64 * )
+		void ktime_get_coarse_real_ts64( struct timespec64 * )
+		void ktime_get_coarse_clocktai_ts64( struct timespec64 * )
+
+	These are quicker than the non-coarse versions, but less accurate,
+	corresponding to CLOCK_MONOTONIC_COARSE and CLOCK_REALTIME_COARSE
+	in user space, along with the equivalent boottime/tai/raw
+	timebase not available in user space.
+
+	The time returned here corresponds to the last timer tick, which
+	may be as much as 10ms in the past (for CONFIG_HZ=100), same as
+	reading the 'jiffies' variable.  These are only useful when called
+	in a fast path and one still expects better than second accuracy,
+	but can't easily use 'jiffies', e.g. for inode timestamps.
+	Skipping the hardware clock access saves around 100 CPU cycles
+	on most modern machines with a reliable cycle counter, but
+	up to several microseconds on older hardware with an external
+	clocksource.
+
+.. c:function:: u64 ktime_get_mono_fast_ns( void )
+		u64 ktime_get_raw_fast_ns( void )
+		u64 ktime_get_boot_fast_ns( void )
+		u64 ktime_get_real_fast_ns( void )
+
+	These variants are safe to call from any context, including from
+	a non-maskable interrupt (NMI) during a timekeeper update, and
+	while we are entering suspend with the clocksource powered down.
+	This is useful in some tracing or debugging code as well as
+	machine check reporting, but most drivers should never call them,
+	since the time is allowed to jump under certain conditions.
+
+Deprecated time interfaces
+--------------------------
+
+Older kernels used some other interfaces that are now being phased out
+but may appear in third-party drivers being ported here. In particular,
+all interfaces returning a 'struct timeval' or 'struct timespec' have
+been replaced because the tv_sec member overflows in year 2038 on 32-bit
+architectures. These are the recommended replacements:
+
+.. c:function:: void ktime_get_ts( struct timespec * )
+
+	Use ktime_get() or ktime_get_ts64() instead.
+
+.. c:function:: void do_gettimeofday( struct timeval * )
+		void getnstimeofday( struct timespec * )
+		void getnstimeofday64( struct timespec64 * )
+		void ktime_get_real_ts( struct timespec * )
+
+	ktime_get_real_ts64() is a direct replacement, but consider using
+	monotonic time (ktime_get_ts64()) and/or a ktime_t based interface
+	(ktime_get()/ktime_get_real()).
+
+.. c:function:: struct timespec current_kernel_time( void )
+		struct timespec64 current_kernel_time64( void )
+		struct timespec get_monotonic_coarse( void )
+		struct timespec64 get_monotonic_coarse64( void )
+
+	These are replaced by ktime_get_coarse_real_ts64() and
+	ktime_get_coarse_ts64(). However, A lot of code that wants
+	coarse-grained times can use the simple 'jiffies' instead, while
+	some drivers may actually want the higher resolution accessors
+	these days.
+
+.. c:function:: struct timespec getrawmonotonic( void )
+		struct timespec64 getrawmonotonic64( void )
+		struct timespec timekeeping_clocktai( void )
+		struct timespec64 timekeeping_clocktai64( void )
+		struct timespec get_monotonic_boottime( void )
+		struct timespec64 get_monotonic_boottime64( void )
+
+	These are replaced by ktime_get_raw()/ktime_get_raw_ts64(),
+	ktime_get_clocktai()/ktime_get_clocktai_ts64() as well
+	as ktime_get_boottime()/ktime_get_boottime_ts64().
+	However, if the particular choice of clock source is not
+	important for the user, consider converting to
+	ktime_get()/ktime_get_ts64() instead for consistency.
diff --git a/Documentation/core-api/tracepoint.rst b/Documentation/core-api/tracepoint.rst
new file mode 100644
index 000000000..6b44bec0d
--- /dev/null
+++ b/Documentation/core-api/tracepoint.rst
@@ -0,0 +1,55 @@
+===============================
+The Linux Kernel Tracepoint API
+===============================
+
+:Author: Jason Baron
+:Author: William Cohen
+
+Introduction
+============
+
+Tracepoints are static probe points that are located in strategic points
+throughout the kernel. 'Probes' register/unregister with tracepoints via
+a callback mechanism. The 'probes' are strictly typed functions that are
+passed a unique set of parameters defined by each tracepoint.
+
+From this simple callback mechanism, 'probes' can be used to profile,
+debug, and understand kernel behavior. There are a number of tools that
+provide a framework for using 'probes'. These tools include Systemtap,
+ftrace, and LTTng.
+
+Tracepoints are defined in a number of header files via various macros.
+Thus, the purpose of this document is to provide a clear accounting of
+the available tracepoints. The intention is to understand not only what
+tracepoints are available but also to understand where future
+tracepoints might be added.
+
+The API presented has functions of the form:
+``trace_tracepointname(function parameters)``. These are the tracepoints
+callbacks that are found throughout the code. Registering and
+unregistering probes with these callback sites is covered in the
+``Documentation/trace/*`` directory.
+
+IRQ
+===
+
+.. kernel-doc:: include/trace/events/irq.h
+   :internal:
+
+SIGNAL
+======
+
+.. kernel-doc:: include/trace/events/signal.h
+   :internal:
+
+Block IO
+========
+
+.. kernel-doc:: include/trace/events/block.h
+   :internal:
+
+Workqueue
+=========
+
+.. kernel-doc:: include/trace/events/workqueue.h
+   :internal:
diff --git a/Documentation/core-api/unaligned-memory-access.rst b/Documentation/core-api/unaligned-memory-access.rst
new file mode 100644
index 000000000..1ee82419d
--- /dev/null
+++ b/Documentation/core-api/unaligned-memory-access.rst
@@ -0,0 +1,265 @@
+=========================
+Unaligned Memory Accesses
+=========================
+
+:Author: Daniel Drake <dsd@gentoo.org>,
+:Author: Johannes Berg <johannes@sipsolutions.net>
+
+:With help from: Alan Cox, Avuton Olrich, Heikki Orsila, Jan Engelhardt,
+  Kyle McMartin, Kyle Moffett, Randy Dunlap, Robert Hancock, Uli Kunitz,
+  Vadim Lobanov
+
+
+Linux runs on a wide variety of architectures which have varying behaviour
+when it comes to memory access. This document presents some details about
+unaligned accesses, why you need to write code that doesn't cause them,
+and how to write such code!
+
+
+The definition of an unaligned access
+=====================================
+
+Unaligned memory accesses occur when you try to read N bytes of data starting
+from an address that is not evenly divisible by N (i.e. addr % N != 0).
+For example, reading 4 bytes of data from address 0x10004 is fine, but
+reading 4 bytes of data from address 0x10005 would be an unaligned memory
+access.
+
+The above may seem a little vague, as memory access can happen in different
+ways. The context here is at the machine code level: certain instructions read
+or write a number of bytes to or from memory (e.g. movb, movw, movl in x86
+assembly). As will become clear, it is relatively easy to spot C statements
+which will compile to multiple-byte memory access instructions, namely when
+dealing with types such as u16, u32 and u64.
+
+
+Natural alignment
+=================
+
+The rule mentioned above forms what we refer to as natural alignment:
+When accessing N bytes of memory, the base memory address must be evenly
+divisible by N, i.e. addr % N == 0.
+
+When writing code, assume the target architecture has natural alignment
+requirements.
+
+In reality, only a few architectures require natural alignment on all sizes
+of memory access. However, we must consider ALL supported architectures;
+writing code that satisfies natural alignment requirements is the easiest way
+to achieve full portability.
+
+
+Why unaligned access is bad
+===========================
+
+The effects of performing an unaligned memory access vary from architecture
+to architecture. It would be easy to write a whole document on the differences
+here; a summary of the common scenarios is presented below:
+
+ - Some architectures are able to perform unaligned memory accesses
+   transparently, but there is usually a significant performance cost.
+ - Some architectures raise processor exceptions when unaligned accesses
+   happen. The exception handler is able to correct the unaligned access,
+   at significant cost to performance.
+ - Some architectures raise processor exceptions when unaligned accesses
+   happen, but the exceptions do not contain enough information for the
+   unaligned access to be corrected.
+ - Some architectures are not capable of unaligned memory access, but will
+   silently perform a different memory access to the one that was requested,
+   resulting in a subtle code bug that is hard to detect!
+
+It should be obvious from the above that if your code causes unaligned
+memory accesses to happen, your code will not work correctly on certain
+platforms and will cause performance problems on others.
+
+
+Code that does not cause unaligned access
+=========================================
+
+At first, the concepts above may seem a little hard to relate to actual
+coding practice. After all, you don't have a great deal of control over
+memory addresses of certain variables, etc.
+
+Fortunately things are not too complex, as in most cases, the compiler
+ensures that things will work for you. For example, take the following
+structure::
+
+	struct foo {
+		u16 field1;
+		u32 field2;
+		u8 field3;
+	};
+
+Let us assume that an instance of the above structure resides in memory
+starting at address 0x10000. With a basic level of understanding, it would
+not be unreasonable to expect that accessing field2 would cause an unaligned
+access. You'd be expecting field2 to be located at offset 2 bytes into the
+structure, i.e. address 0x10002, but that address is not evenly divisible
+by 4 (remember, we're reading a 4 byte value here).
+
+Fortunately, the compiler understands the alignment constraints, so in the
+above case it would insert 2 bytes of padding in between field1 and field2.
+Therefore, for standard structure types you can always rely on the compiler
+to pad structures so that accesses to fields are suitably aligned (assuming
+you do not cast the field to a type of different length).
+
+Similarly, you can also rely on the compiler to align variables and function
+parameters to a naturally aligned scheme, based on the size of the type of
+the variable.
+
+At this point, it should be clear that accessing a single byte (u8 or char)
+will never cause an unaligned access, because all memory addresses are evenly
+divisible by one.
+
+On a related topic, with the above considerations in mind you may observe
+that you could reorder the fields in the structure in order to place fields
+where padding would otherwise be inserted, and hence reduce the overall
+resident memory size of structure instances. The optimal layout of the
+above example is::
+
+	struct foo {
+		u32 field2;
+		u16 field1;
+		u8 field3;
+	};
+
+For a natural alignment scheme, the compiler would only have to add a single
+byte of padding at the end of the structure. This padding is added in order
+to satisfy alignment constraints for arrays of these structures.
+
+Another point worth mentioning is the use of __attribute__((packed)) on a
+structure type. This GCC-specific attribute tells the compiler never to
+insert any padding within structures, useful when you want to use a C struct
+to represent some data that comes in a fixed arrangement 'off the wire'.
+
+You might be inclined to believe that usage of this attribute can easily
+lead to unaligned accesses when accessing fields that do not satisfy
+architectural alignment requirements. However, again, the compiler is aware
+of the alignment constraints and will generate extra instructions to perform
+the memory access in a way that does not cause unaligned access. Of course,
+the extra instructions obviously cause a loss in performance compared to the
+non-packed case, so the packed attribute should only be used when avoiding
+structure padding is of importance.
+
+
+Code that causes unaligned access
+=================================
+
+With the above in mind, let's move onto a real life example of a function
+that can cause an unaligned memory access. The following function taken
+from include/linux/etherdevice.h is an optimized routine to compare two
+ethernet MAC addresses for equality::
+
+  bool ether_addr_equal(const u8 *addr1, const u8 *addr2)
+  {
+  #ifdef CONFIG_HAVE_EFFICIENT_UNALIGNED_ACCESS
+	u32 fold = ((*(const u32 *)addr1) ^ (*(const u32 *)addr2)) |
+		   ((*(const u16 *)(addr1 + 4)) ^ (*(const u16 *)(addr2 + 4)));
+
+	return fold == 0;
+  #else
+	const u16 *a = (const u16 *)addr1;
+	const u16 *b = (const u16 *)addr2;
+	return ((a[0] ^ b[0]) | (a[1] ^ b[1]) | (a[2] ^ b[2])) == 0;
+  #endif
+  }
+
+In the above function, when the hardware has efficient unaligned access
+capability, there is no issue with this code.  But when the hardware isn't
+able to access memory on arbitrary boundaries, the reference to a[0] causes
+2 bytes (16 bits) to be read from memory starting at address addr1.
+
+Think about what would happen if addr1 was an odd address such as 0x10003.
+(Hint: it'd be an unaligned access.)
+
+Despite the potential unaligned access problems with the above function, it
+is included in the kernel anyway but is understood to only work normally on
+16-bit-aligned addresses. It is up to the caller to ensure this alignment or
+not use this function at all. This alignment-unsafe function is still useful
+as it is a decent optimization for the cases when you can ensure alignment,
+which is true almost all of the time in ethernet networking context.
+
+
+Here is another example of some code that could cause unaligned accesses::
+
+	void myfunc(u8 *data, u32 value)
+	{
+		[...]
+		*((u32 *) data) = cpu_to_le32(value);
+		[...]
+	}
+
+This code will cause unaligned accesses every time the data parameter points
+to an address that is not evenly divisible by 4.
+
+In summary, the 2 main scenarios where you may run into unaligned access
+problems involve:
+
+ 1. Casting variables to types of different lengths
+ 2. Pointer arithmetic followed by access to at least 2 bytes of data
+
+
+Avoiding unaligned accesses
+===========================
+
+The easiest way to avoid unaligned access is to use the get_unaligned() and
+put_unaligned() macros provided by the <asm/unaligned.h> header file.
+
+Going back to an earlier example of code that potentially causes unaligned
+access::
+
+	void myfunc(u8 *data, u32 value)
+	{
+		[...]
+		*((u32 *) data) = cpu_to_le32(value);
+		[...]
+	}
+
+To avoid the unaligned memory access, you would rewrite it as follows::
+
+	void myfunc(u8 *data, u32 value)
+	{
+		[...]
+		value = cpu_to_le32(value);
+		put_unaligned(value, (u32 *) data);
+		[...]
+	}
+
+The get_unaligned() macro works similarly. Assuming 'data' is a pointer to
+memory and you wish to avoid unaligned access, its usage is as follows::
+
+	u32 value = get_unaligned((u32 *) data);
+
+These macros work for memory accesses of any length (not just 32 bits as
+in the examples above). Be aware that when compared to standard access of
+aligned memory, using these macros to access unaligned memory can be costly in
+terms of performance.
+
+If use of such macros is not convenient, another option is to use memcpy(),
+where the source or destination (or both) are of type u8* or unsigned char*.
+Due to the byte-wise nature of this operation, unaligned accesses are avoided.
+
+
+Alignment vs. Networking
+========================
+
+On architectures that require aligned loads, networking requires that the IP
+header is aligned on a four-byte boundary to optimise the IP stack. For
+regular ethernet hardware, the constant NET_IP_ALIGN is used. On most
+architectures this constant has the value 2 because the normal ethernet
+header is 14 bytes long, so in order to get proper alignment one needs to
+DMA to an address which can be expressed as 4*n + 2. One notable exception
+here is powerpc which defines NET_IP_ALIGN to 0 because DMA to unaligned
+addresses can be very expensive and dwarf the cost of unaligned loads.
+
+For some ethernet hardware that cannot DMA to unaligned addresses like
+4*n+2 or non-ethernet hardware, this can be a problem, and it is then
+required to copy the incoming frame into an aligned buffer. Because this is
+unnecessary on architectures that can do unaligned accesses, the code can be
+made dependent on CONFIG_HAVE_EFFICIENT_UNALIGNED_ACCESS like so::
+
+	#ifdef CONFIG_HAVE_EFFICIENT_UNALIGNED_ACCESS
+		skb = original skb
+	#else
+		skb = copy skb
+	#endif
diff --git a/Documentation/core-api/workqueue.rst b/Documentation/core-api/workqueue.rst
new file mode 100644
index 000000000..541d31de8
--- /dev/null
+++ b/Documentation/core-api/workqueue.rst
@@ -0,0 +1,400 @@
+====================================
+Concurrency Managed Workqueue (cmwq)
+====================================
+
+:Date: September, 2010
+:Author: Tejun Heo <tj@kernel.org>
+:Author: Florian Mickler <florian@mickler.org>
+
+
+Introduction
+============
+
+There are many cases where an asynchronous process execution context
+is needed and the workqueue (wq) API is the most commonly used
+mechanism for such cases.
+
+When such an asynchronous execution context is needed, a work item
+describing which function to execute is put on a queue.  An
+independent thread serves as the asynchronous execution context.  The
+queue is called workqueue and the thread is called worker.
+
+While there are work items on the workqueue the worker executes the
+functions associated with the work items one after the other.  When
+there is no work item left on the workqueue the worker becomes idle.
+When a new work item gets queued, the worker begins executing again.
+
+
+Why cmwq?
+=========
+
+In the original wq implementation, a multi threaded (MT) wq had one
+worker thread per CPU and a single threaded (ST) wq had one worker
+thread system-wide.  A single MT wq needed to keep around the same
+number of workers as the number of CPUs.  The kernel grew a lot of MT
+wq users over the years and with the number of CPU cores continuously
+rising, some systems saturated the default 32k PID space just booting
+up.
+
+Although MT wq wasted a lot of resource, the level of concurrency
+provided was unsatisfactory.  The limitation was common to both ST and
+MT wq albeit less severe on MT.  Each wq maintained its own separate
+worker pool.  An MT wq could provide only one execution context per CPU
+while an ST wq one for the whole system.  Work items had to compete for
+those very limited execution contexts leading to various problems
+including proneness to deadlocks around the single execution context.
+
+The tension between the provided level of concurrency and resource
+usage also forced its users to make unnecessary tradeoffs like libata
+choosing to use ST wq for polling PIOs and accepting an unnecessary
+limitation that no two polling PIOs can progress at the same time.  As
+MT wq don't provide much better concurrency, users which require
+higher level of concurrency, like async or fscache, had to implement
+their own thread pool.
+
+Concurrency Managed Workqueue (cmwq) is a reimplementation of wq with
+focus on the following goals.
+
+* Maintain compatibility with the original workqueue API.
+
+* Use per-CPU unified worker pools shared by all wq to provide
+  flexible level of concurrency on demand without wasting a lot of
+  resource.
+
+* Automatically regulate worker pool and level of concurrency so that
+  the API users don't need to worry about such details.
+
+
+The Design
+==========
+
+In order to ease the asynchronous execution of functions a new
+abstraction, the work item, is introduced.
+
+A work item is a simple struct that holds a pointer to the function
+that is to be executed asynchronously.  Whenever a driver or subsystem
+wants a function to be executed asynchronously it has to set up a work
+item pointing to that function and queue that work item on a
+workqueue.
+
+Special purpose threads, called worker threads, execute the functions
+off of the queue, one after the other.  If no work is queued, the
+worker threads become idle.  These worker threads are managed in so
+called worker-pools.
+
+The cmwq design differentiates between the user-facing workqueues that
+subsystems and drivers queue work items on and the backend mechanism
+which manages worker-pools and processes the queued work items.
+
+There are two worker-pools, one for normal work items and the other
+for high priority ones, for each possible CPU and some extra
+worker-pools to serve work items queued on unbound workqueues - the
+number of these backing pools is dynamic.
+
+Subsystems and drivers can create and queue work items through special
+workqueue API functions as they see fit. They can influence some
+aspects of the way the work items are executed by setting flags on the
+workqueue they are putting the work item on. These flags include
+things like CPU locality, concurrency limits, priority and more.  To
+get a detailed overview refer to the API description of
+``alloc_workqueue()`` below.
+
+When a work item is queued to a workqueue, the target worker-pool is
+determined according to the queue parameters and workqueue attributes
+and appended on the shared worklist of the worker-pool.  For example,
+unless specifically overridden, a work item of a bound workqueue will
+be queued on the worklist of either normal or highpri worker-pool that
+is associated to the CPU the issuer is running on.
+
+For any worker pool implementation, managing the concurrency level
+(how many execution contexts are active) is an important issue.  cmwq
+tries to keep the concurrency at a minimal but sufficient level.
+Minimal to save resources and sufficient in that the system is used at
+its full capacity.
+
+Each worker-pool bound to an actual CPU implements concurrency
+management by hooking into the scheduler.  The worker-pool is notified
+whenever an active worker wakes up or sleeps and keeps track of the
+number of the currently runnable workers.  Generally, work items are
+not expected to hog a CPU and consume many cycles.  That means
+maintaining just enough concurrency to prevent work processing from
+stalling should be optimal.  As long as there are one or more runnable
+workers on the CPU, the worker-pool doesn't start execution of a new
+work, but, when the last running worker goes to sleep, it immediately
+schedules a new worker so that the CPU doesn't sit idle while there
+are pending work items.  This allows using a minimal number of workers
+without losing execution bandwidth.
+
+Keeping idle workers around doesn't cost other than the memory space
+for kthreads, so cmwq holds onto idle ones for a while before killing
+them.
+
+For unbound workqueues, the number of backing pools is dynamic.
+Unbound workqueue can be assigned custom attributes using
+``apply_workqueue_attrs()`` and workqueue will automatically create
+backing worker pools matching the attributes.  The responsibility of
+regulating concurrency level is on the users.  There is also a flag to
+mark a bound wq to ignore the concurrency management.  Please refer to
+the API section for details.
+
+Forward progress guarantee relies on that workers can be created when
+more execution contexts are necessary, which in turn is guaranteed
+through the use of rescue workers.  All work items which might be used
+on code paths that handle memory reclaim are required to be queued on
+wq's that have a rescue-worker reserved for execution under memory
+pressure.  Else it is possible that the worker-pool deadlocks waiting
+for execution contexts to free up.
+
+
+Application Programming Interface (API)
+=======================================
+
+``alloc_workqueue()`` allocates a wq.  The original
+``create_*workqueue()`` functions are deprecated and scheduled for
+removal.  ``alloc_workqueue()`` takes three arguments - ``@name``,
+``@flags`` and ``@max_active``.  ``@name`` is the name of the wq and
+also used as the name of the rescuer thread if there is one.
+
+A wq no longer manages execution resources but serves as a domain for
+forward progress guarantee, flush and work item attributes. ``@flags``
+and ``@max_active`` control how work items are assigned execution
+resources, scheduled and executed.
+
+
+``flags``
+---------
+
+``WQ_UNBOUND``
+  Work items queued to an unbound wq are served by the special
+  worker-pools which host workers which are not bound to any
+  specific CPU.  This makes the wq behave as a simple execution
+  context provider without concurrency management.  The unbound
+  worker-pools try to start execution of work items as soon as
+  possible.  Unbound wq sacrifices locality but is useful for
+  the following cases.
+
+  * Wide fluctuation in the concurrency level requirement is
+    expected and using bound wq may end up creating large number
+    of mostly unused workers across different CPUs as the issuer
+    hops through different CPUs.
+
+  * Long running CPU intensive workloads which can be better
+    managed by the system scheduler.
+
+``WQ_FREEZABLE``
+  A freezable wq participates in the freeze phase of the system
+  suspend operations.  Work items on the wq are drained and no
+  new work item starts execution until thawed.
+
+``WQ_MEM_RECLAIM``
+  All wq which might be used in the memory reclaim paths **MUST**
+  have this flag set.  The wq is guaranteed to have at least one
+  execution context regardless of memory pressure.
+
+``WQ_HIGHPRI``
+  Work items of a highpri wq are queued to the highpri
+  worker-pool of the target cpu.  Highpri worker-pools are
+  served by worker threads with elevated nice level.
+
+  Note that normal and highpri worker-pools don't interact with
+  each other.  Each maintains its separate pool of workers and
+  implements concurrency management among its workers.
+
+``WQ_CPU_INTENSIVE``
+  Work items of a CPU intensive wq do not contribute to the
+  concurrency level.  In other words, runnable CPU intensive
+  work items will not prevent other work items in the same
+  worker-pool from starting execution.  This is useful for bound
+  work items which are expected to hog CPU cycles so that their
+  execution is regulated by the system scheduler.
+
+  Although CPU intensive work items don't contribute to the
+  concurrency level, start of their executions is still
+  regulated by the concurrency management and runnable
+  non-CPU-intensive work items can delay execution of CPU
+  intensive work items.
+
+  This flag is meaningless for unbound wq.
+
+Note that the flag ``WQ_NON_REENTRANT`` no longer exists as all
+workqueues are now non-reentrant - any work item is guaranteed to be
+executed by at most one worker system-wide at any given time.
+
+
+``max_active``
+--------------
+
+``@max_active`` determines the maximum number of execution contexts
+per CPU which can be assigned to the work items of a wq.  For example,
+with ``@max_active`` of 16, at most 16 work items of the wq can be
+executing at the same time per CPU.
+
+Currently, for a bound wq, the maximum limit for ``@max_active`` is
+512 and the default value used when 0 is specified is 256.  For an
+unbound wq, the limit is higher of 512 and 4 *
+``num_possible_cpus()``.  These values are chosen sufficiently high
+such that they are not the limiting factor while providing protection
+in runaway cases.
+
+The number of active work items of a wq is usually regulated by the
+users of the wq, more specifically, by how many work items the users
+may queue at the same time.  Unless there is a specific need for
+throttling the number of active work items, specifying '0' is
+recommended.
+
+Some users depend on the strict execution ordering of ST wq.  The
+combination of ``@max_active`` of 1 and ``WQ_UNBOUND`` used to
+achieve this behavior.  Work items on such wq were always queued to the
+unbound worker-pools and only one work item could be active at any given
+time thus achieving the same ordering property as ST wq.
+
+In the current implementation the above configuration only guarantees
+ST behavior within a given NUMA node. Instead ``alloc_ordered_queue()`` should
+be used to achieve system-wide ST behavior.
+
+
+Example Execution Scenarios
+===========================
+
+The following example execution scenarios try to illustrate how cmwq
+behave under different configurations.
+
+ Work items w0, w1, w2 are queued to a bound wq q0 on the same CPU.
+ w0 burns CPU for 5ms then sleeps for 10ms then burns CPU for 5ms
+ again before finishing.  w1 and w2 burn CPU for 5ms then sleep for
+ 10ms.
+
+Ignoring all other tasks, works and processing overhead, and assuming
+simple FIFO scheduling, the following is one highly simplified version
+of possible sequences of events with the original wq. ::
+
+ TIME IN MSECS	EVENT
+ 0		w0 starts and burns CPU
+ 5		w0 sleeps
+ 15		w0 wakes up and burns CPU
+ 20		w0 finishes
+ 20		w1 starts and burns CPU
+ 25		w1 sleeps
+ 35		w1 wakes up and finishes
+ 35		w2 starts and burns CPU
+ 40		w2 sleeps
+ 50		w2 wakes up and finishes
+
+And with cmwq with ``@max_active`` >= 3, ::
+
+ TIME IN MSECS	EVENT
+ 0		w0 starts and burns CPU
+ 5		w0 sleeps
+ 5		w1 starts and burns CPU
+ 10		w1 sleeps
+ 10		w2 starts and burns CPU
+ 15		w2 sleeps
+ 15		w0 wakes up and burns CPU
+ 20		w0 finishes
+ 20		w1 wakes up and finishes
+ 25		w2 wakes up and finishes
+
+If ``@max_active`` == 2, ::
+
+ TIME IN MSECS	EVENT
+ 0		w0 starts and burns CPU
+ 5		w0 sleeps
+ 5		w1 starts and burns CPU
+ 10		w1 sleeps
+ 15		w0 wakes up and burns CPU
+ 20		w0 finishes
+ 20		w1 wakes up and finishes
+ 20		w2 starts and burns CPU
+ 25		w2 sleeps
+ 35		w2 wakes up and finishes
+
+Now, let's assume w1 and w2 are queued to a different wq q1 which has
+``WQ_CPU_INTENSIVE`` set, ::
+
+ TIME IN MSECS	EVENT
+ 0		w0 starts and burns CPU
+ 5		w0 sleeps
+ 5		w1 and w2 start and burn CPU
+ 10		w1 sleeps
+ 15		w2 sleeps
+ 15		w0 wakes up and burns CPU
+ 20		w0 finishes
+ 20		w1 wakes up and finishes
+ 25		w2 wakes up and finishes
+
+
+Guidelines
+==========
+
+* Do not forget to use ``WQ_MEM_RECLAIM`` if a wq may process work
+  items which are used during memory reclaim.  Each wq with
+  ``WQ_MEM_RECLAIM`` set has an execution context reserved for it.  If
+  there is dependency among multiple work items used during memory
+  reclaim, they should be queued to separate wq each with
+  ``WQ_MEM_RECLAIM``.
+
+* Unless strict ordering is required, there is no need to use ST wq.
+
+* Unless there is a specific need, using 0 for @max_active is
+  recommended.  In most use cases, concurrency level usually stays
+  well under the default limit.
+
+* A wq serves as a domain for forward progress guarantee
+  (``WQ_MEM_RECLAIM``, flush and work item attributes.  Work items
+  which are not involved in memory reclaim and don't need to be
+  flushed as a part of a group of work items, and don't require any
+  special attribute, can use one of the system wq.  There is no
+  difference in execution characteristics between using a dedicated wq
+  and a system wq.
+
+* Unless work items are expected to consume a huge amount of CPU
+  cycles, using a bound wq is usually beneficial due to the increased
+  level of locality in wq operations and work item execution.
+
+
+Debugging
+=========
+
+Because the work functions are executed by generic worker threads
+there are a few tricks needed to shed some light on misbehaving
+workqueue users.
+
+Worker threads show up in the process list as: ::
+
+  root      5671  0.0  0.0      0     0 ?        S    12:07   0:00 [kworker/0:1]
+  root      5672  0.0  0.0      0     0 ?        S    12:07   0:00 [kworker/1:2]
+  root      5673  0.0  0.0      0     0 ?        S    12:12   0:00 [kworker/0:0]
+  root      5674  0.0  0.0      0     0 ?        S    12:13   0:00 [kworker/1:0]
+
+If kworkers are going crazy (using too much cpu), there are two types
+of possible problems:
+
+	1. Something being scheduled in rapid succession
+	2. A single work item that consumes lots of cpu cycles
+
+The first one can be tracked using tracing: ::
+
+	$ echo workqueue:workqueue_queue_work > /sys/kernel/debug/tracing/set_event
+	$ cat /sys/kernel/debug/tracing/trace_pipe > out.txt
+	(wait a few secs)
+	^C
+
+If something is busy looping on work queueing, it would be dominating
+the output and the offender can be determined with the work item
+function.
+
+For the second type of problems it should be possible to just check
+the stack trace of the offending worker thread. ::
+
+	$ cat /proc/THE_OFFENDING_KWORKER/stack
+
+The work item's function should be trivially visible in the stack
+trace.
+
+
+Kernel Inline Documentations Reference
+======================================
+
+.. kernel-doc:: include/linux/workqueue.h
+
+.. kernel-doc:: kernel/workqueue.c
diff --git a/Documentation/core-api/xarray.rst b/Documentation/core-api/xarray.rst
new file mode 100644
index 000000000..a137a0e6d
--- /dev/null
+++ b/Documentation/core-api/xarray.rst
@@ -0,0 +1,492 @@
+.. SPDX-License-Identifier: GPL-2.0+
+
+======
+XArray
+======
+
+:Author: Matthew Wilcox
+
+Overview
+========
+
+The XArray is an abstract data type which behaves like a very large array
+of pointers.  It meets many of the same needs as a hash or a conventional
+resizable array.  Unlike a hash, it allows you to sensibly go to the
+next or previous entry in a cache-efficient manner.  In contrast to a
+resizable array, there is no need to copy data or change MMU mappings in
+order to grow the array.  It is more memory-efficient, parallelisable
+and cache friendly than a doubly-linked list.  It takes advantage of
+RCU to perform lookups without locking.
+
+The XArray implementation is efficient when the indices used are densely
+clustered; hashing the object and using the hash as the index will not
+perform well.  The XArray is optimised for small indices, but still has
+good performance with large indices.  If your index can be larger than
+``ULONG_MAX`` then the XArray is not the data type for you.  The most
+important user of the XArray is the page cache.
+
+Normal pointers may be stored in the XArray directly.  They must be 4-byte
+aligned, which is true for any pointer returned from kmalloc() and
+alloc_page().  It isn't true for arbitrary user-space pointers,
+nor for function pointers.  You can store pointers to statically allocated
+objects, as long as those objects have an alignment of at least 4.
+
+You can also store integers between 0 and ``LONG_MAX`` in the XArray.
+You must first convert it into an entry using xa_mk_value().
+When you retrieve an entry from the XArray, you can check whether it is
+a value entry by calling xa_is_value(), and convert it back to
+an integer by calling xa_to_value().
+
+Some users want to tag the pointers they store in the XArray.  You can
+call xa_tag_pointer() to create an entry with a tag, xa_untag_pointer()
+to turn a tagged entry back into an untagged pointer and xa_pointer_tag()
+to retrieve the tag of an entry.  Tagged pointers use the same bits that
+are used to distinguish value entries from normal pointers, so you must
+decide whether they want to store value entries or tagged pointers in
+any particular XArray.
+
+The XArray does not support storing IS_ERR() pointers as some
+conflict with value entries or internal entries.
+
+An unusual feature of the XArray is the ability to create entries which
+occupy a range of indices.  Once stored to, looking up any index in
+the range will return the same entry as looking up any other index in
+the range.  Storing to any index will store to all of them.  Multi-index
+entries can be explicitly split into smaller entries, or storing ``NULL``
+into any entry will cause the XArray to forget about the range.
+
+Normal API
+==========
+
+Start by initialising an XArray, either with DEFINE_XARRAY()
+for statically allocated XArrays or xa_init() for dynamically
+allocated ones.  A freshly-initialised XArray contains a ``NULL``
+pointer at every index.
+
+You can then set entries using xa_store() and get entries
+using xa_load().  xa_store will overwrite any entry with the
+new entry and return the previous entry stored at that index.  You can
+use xa_erase() instead of calling xa_store() with a
+``NULL`` entry.  There is no difference between an entry that has never
+been stored to, one that has been erased and one that has most recently
+had ``NULL`` stored to it.
+
+You can conditionally replace an entry at an index by using
+xa_cmpxchg().  Like cmpxchg(), it will only succeed if
+the entry at that index has the 'old' value.  It also returns the entry
+which was at that index; if it returns the same entry which was passed as
+'old', then xa_cmpxchg() succeeded.
+
+If you want to only store a new entry to an index if the current entry
+at that index is ``NULL``, you can use xa_insert() which
+returns ``-EBUSY`` if the entry is not empty.
+
+You can copy entries out of the XArray into a plain array by calling
+xa_extract().  Or you can iterate over the present entries in the XArray
+by calling xa_for_each(), xa_for_each_start() or xa_for_each_range().
+You may prefer to use xa_find() or xa_find_after() to move to the next
+present entry in the XArray.
+
+Calling xa_store_range() stores the same entry in a range
+of indices.  If you do this, some of the other operations will behave
+in a slightly odd way.  For example, marking the entry at one index
+may result in the entry being marked at some, but not all of the other
+indices.  Storing into one index may result in the entry retrieved by
+some, but not all of the other indices changing.
+
+Sometimes you need to ensure that a subsequent call to xa_store()
+will not need to allocate memory.  The xa_reserve() function
+will store a reserved entry at the indicated index.  Users of the
+normal API will see this entry as containing ``NULL``.  If you do
+not need to use the reserved entry, you can call xa_release()
+to remove the unused entry.  If another user has stored to the entry
+in the meantime, xa_release() will do nothing; if instead you
+want the entry to become ``NULL``, you should use xa_erase().
+Using xa_insert() on a reserved entry will fail.
+
+If all entries in the array are ``NULL``, the xa_empty() function
+will return ``true``.
+
+Finally, you can remove all entries from an XArray by calling
+xa_destroy().  If the XArray entries are pointers, you may wish
+to free the entries first.  You can do this by iterating over all present
+entries in the XArray using the xa_for_each() iterator.
+
+Search Marks
+------------
+
+Each entry in the array has three bits associated with it called marks.
+Each mark may be set or cleared independently of the others.  You can
+iterate over marked entries by using the xa_for_each_marked() iterator.
+
+You can enquire whether a mark is set on an entry by using
+xa_get_mark().  If the entry is not ``NULL``, you can set a mark on it
+by using xa_set_mark() and remove the mark from an entry by calling
+xa_clear_mark().  You can ask whether any entry in the XArray has a
+particular mark set by calling xa_marked().  Erasing an entry from the
+XArray causes all marks associated with that entry to be cleared.
+
+Setting or clearing a mark on any index of a multi-index entry will
+affect all indices covered by that entry.  Querying the mark on any
+index will return the same result.
+
+There is no way to iterate over entries which are not marked; the data
+structure does not allow this to be implemented efficiently.  There are
+not currently iterators to search for logical combinations of bits (eg
+iterate over all entries which have both ``XA_MARK_1`` and ``XA_MARK_2``
+set, or iterate over all entries which have ``XA_MARK_0`` or ``XA_MARK_2``
+set).  It would be possible to add these if a user arises.
+
+Allocating XArrays
+------------------
+
+If you use DEFINE_XARRAY_ALLOC() to define the XArray, or
+initialise it by passing ``XA_FLAGS_ALLOC`` to xa_init_flags(),
+the XArray changes to track whether entries are in use or not.
+
+You can call xa_alloc() to store the entry at an unused index
+in the XArray.  If you need to modify the array from interrupt context,
+you can use xa_alloc_bh() or xa_alloc_irq() to disable
+interrupts while allocating the ID.
+
+Using xa_store(), xa_cmpxchg() or xa_insert() will
+also mark the entry as being allocated.  Unlike a normal XArray, storing
+``NULL`` will mark the entry as being in use, like xa_reserve().
+To free an entry, use xa_erase() (or xa_release() if
+you only want to free the entry if it's ``NULL``).
+
+By default, the lowest free entry is allocated starting from 0.  If you
+want to allocate entries starting at 1, it is more efficient to use
+DEFINE_XARRAY_ALLOC1() or ``XA_FLAGS_ALLOC1``.  If you want to
+allocate IDs up to a maximum, then wrap back around to the lowest free
+ID, you can use xa_alloc_cyclic().
+
+You cannot use ``XA_MARK_0`` with an allocating XArray as this mark
+is used to track whether an entry is free or not.  The other marks are
+available for your use.
+
+Memory allocation
+-----------------
+
+The xa_store(), xa_cmpxchg(), xa_alloc(),
+xa_reserve() and xa_insert() functions take a gfp_t
+parameter in case the XArray needs to allocate memory to store this entry.
+If the entry is being deleted, no memory allocation needs to be performed,
+and the GFP flags specified will be ignored.
+
+It is possible for no memory to be allocatable, particularly if you pass
+a restrictive set of GFP flags.  In that case, the functions return a
+special value which can be turned into an errno using xa_err().
+If you don't need to know exactly which error occurred, using
+xa_is_err() is slightly more efficient.
+
+Locking
+-------
+
+When using the Normal API, you do not have to worry about locking.
+The XArray uses RCU and an internal spinlock to synchronise access:
+
+No lock needed:
+ * xa_empty()
+ * xa_marked()
+
+Takes RCU read lock:
+ * xa_load()
+ * xa_for_each()
+ * xa_for_each_start()
+ * xa_for_each_range()
+ * xa_find()
+ * xa_find_after()
+ * xa_extract()
+ * xa_get_mark()
+
+Takes xa_lock internally:
+ * xa_store()
+ * xa_store_bh()
+ * xa_store_irq()
+ * xa_insert()
+ * xa_insert_bh()
+ * xa_insert_irq()
+ * xa_erase()
+ * xa_erase_bh()
+ * xa_erase_irq()
+ * xa_cmpxchg()
+ * xa_cmpxchg_bh()
+ * xa_cmpxchg_irq()
+ * xa_store_range()
+ * xa_alloc()
+ * xa_alloc_bh()
+ * xa_alloc_irq()
+ * xa_reserve()
+ * xa_reserve_bh()
+ * xa_reserve_irq()
+ * xa_destroy()
+ * xa_set_mark()
+ * xa_clear_mark()
+
+Assumes xa_lock held on entry:
+ * __xa_store()
+ * __xa_insert()
+ * __xa_erase()
+ * __xa_cmpxchg()
+ * __xa_alloc()
+ * __xa_set_mark()
+ * __xa_clear_mark()
+
+If you want to take advantage of the lock to protect the data structures
+that you are storing in the XArray, you can call xa_lock()
+before calling xa_load(), then take a reference count on the
+object you have found before calling xa_unlock().  This will
+prevent stores from removing the object from the array between looking
+up the object and incrementing the refcount.  You can also use RCU to
+avoid dereferencing freed memory, but an explanation of that is beyond
+the scope of this document.
+
+The XArray does not disable interrupts or softirqs while modifying
+the array.  It is safe to read the XArray from interrupt or softirq
+context as the RCU lock provides enough protection.
+
+If, for example, you want to store entries in the XArray in process
+context and then erase them in softirq context, you can do that this way::
+
+    void foo_init(struct foo *foo)
+    {
+        xa_init_flags(&foo->array, XA_FLAGS_LOCK_BH);
+    }
+
+    int foo_store(struct foo *foo, unsigned long index, void *entry)
+    {
+        int err;
+
+        xa_lock_bh(&foo->array);
+        err = xa_err(__xa_store(&foo->array, index, entry, GFP_KERNEL));
+        if (!err)
+            foo->count++;
+        xa_unlock_bh(&foo->array);
+        return err;
+    }
+
+    /* foo_erase() is only called from softirq context */
+    void foo_erase(struct foo *foo, unsigned long index)
+    {
+        xa_lock(&foo->array);
+        __xa_erase(&foo->array, index);
+        foo->count--;
+        xa_unlock(&foo->array);
+    }
+
+If you are going to modify the XArray from interrupt or softirq context,
+you need to initialise the array using xa_init_flags(), passing
+``XA_FLAGS_LOCK_IRQ`` or ``XA_FLAGS_LOCK_BH``.
+
+The above example also shows a common pattern of wanting to extend the
+coverage of the xa_lock on the store side to protect some statistics
+associated with the array.
+
+Sharing the XArray with interrupt context is also possible, either
+using xa_lock_irqsave() in both the interrupt handler and process
+context, or xa_lock_irq() in process context and xa_lock()
+in the interrupt handler.  Some of the more common patterns have helper
+functions such as xa_store_bh(), xa_store_irq(),
+xa_erase_bh(), xa_erase_irq(), xa_cmpxchg_bh()
+and xa_cmpxchg_irq().
+
+Sometimes you need to protect access to the XArray with a mutex because
+that lock sits above another mutex in the locking hierarchy.  That does
+not entitle you to use functions like __xa_erase() without taking
+the xa_lock; the xa_lock is used for lockdep validation and will be used
+for other purposes in the future.
+
+The __xa_set_mark() and __xa_clear_mark() functions are also
+available for situations where you look up an entry and want to atomically
+set or clear a mark.  It may be more efficient to use the advanced API
+in this case, as it will save you from walking the tree twice.
+
+Advanced API
+============
+
+The advanced API offers more flexibility and better performance at the
+cost of an interface which can be harder to use and has fewer safeguards.
+No locking is done for you by the advanced API, and you are required
+to use the xa_lock while modifying the array.  You can choose whether
+to use the xa_lock or the RCU lock while doing read-only operations on
+the array.  You can mix advanced and normal operations on the same array;
+indeed the normal API is implemented in terms of the advanced API.  The
+advanced API is only available to modules with a GPL-compatible license.
+
+The advanced API is based around the xa_state.  This is an opaque data
+structure which you declare on the stack using the XA_STATE()
+macro.  This macro initialises the xa_state ready to start walking
+around the XArray.  It is used as a cursor to maintain the position
+in the XArray and let you compose various operations together without
+having to restart from the top every time.
+
+The xa_state is also used to store errors.  You can call
+xas_error() to retrieve the error.  All operations check whether
+the xa_state is in an error state before proceeding, so there's no need
+for you to check for an error after each call; you can make multiple
+calls in succession and only check at a convenient point.  The only
+errors currently generated by the XArray code itself are ``ENOMEM`` and
+``EINVAL``, but it supports arbitrary errors in case you want to call
+xas_set_err() yourself.
+
+If the xa_state is holding an ``ENOMEM`` error, calling xas_nomem()
+will attempt to allocate more memory using the specified gfp flags and
+cache it in the xa_state for the next attempt.  The idea is that you take
+the xa_lock, attempt the operation and drop the lock.  The operation
+attempts to allocate memory while holding the lock, but it is more
+likely to fail.  Once you have dropped the lock, xas_nomem()
+can try harder to allocate more memory.  It will return ``true`` if it
+is worth retrying the operation (i.e. that there was a memory error *and*
+more memory was allocated).  If it has previously allocated memory, and
+that memory wasn't used, and there is no error (or some error that isn't
+``ENOMEM``), then it will free the memory previously allocated.
+
+Internal Entries
+----------------
+
+The XArray reserves some entries for its own purposes.  These are never
+exposed through the normal API, but when using the advanced API, it's
+possible to see them.  Usually the best way to handle them is to pass them
+to xas_retry(), and retry the operation if it returns ``true``.
+
+.. flat-table::
+   :widths: 1 1 6
+
+   * - Name
+     - Test
+     - Usage
+
+   * - Node
+     - xa_is_node()
+     - An XArray node.  May be visible when using a multi-index xa_state.
+
+   * - Sibling
+     - xa_is_sibling()
+     - A non-canonical entry for a multi-index entry.  The value indicates
+       which slot in this node has the canonical entry.
+
+   * - Retry
+     - xa_is_retry()
+     - This entry is currently being modified by a thread which has the
+       xa_lock.  The node containing this entry may be freed at the end
+       of this RCU period.  You should restart the lookup from the head
+       of the array.
+
+   * - Zero
+     - xa_is_zero()
+     - Zero entries appear as ``NULL`` through the Normal API, but occupy
+       an entry in the XArray which can be used to reserve the index for
+       future use.  This is used by allocating XArrays for allocated entries
+       which are ``NULL``.
+
+Other internal entries may be added in the future.  As far as possible, they
+will be handled by xas_retry().
+
+Additional functionality
+------------------------
+
+The xas_create_range() function allocates all the necessary memory
+to store every entry in a range.  It will set ENOMEM in the xa_state if
+it cannot allocate memory.
+
+You can use xas_init_marks() to reset the marks on an entry
+to their default state.  This is usually all marks clear, unless the
+XArray is marked with ``XA_FLAGS_TRACK_FREE``, in which case mark 0 is set
+and all other marks are clear.  Replacing one entry with another using
+xas_store() will not reset the marks on that entry; if you want
+the marks reset, you should do that explicitly.
+
+The xas_load() will walk the xa_state as close to the entry
+as it can.  If you know the xa_state has already been walked to the
+entry and need to check that the entry hasn't changed, you can use
+xas_reload() to save a function call.
+
+If you need to move to a different index in the XArray, call
+xas_set().  This resets the cursor to the top of the tree, which
+will generally make the next operation walk the cursor to the desired
+spot in the tree.  If you want to move to the next or previous index,
+call xas_next() or xas_prev().  Setting the index does
+not walk the cursor around the array so does not require a lock to be
+held, while moving to the next or previous index does.
+
+You can search for the next present entry using xas_find().  This
+is the equivalent of both xa_find() and xa_find_after();
+if the cursor has been walked to an entry, then it will find the next
+entry after the one currently referenced.  If not, it will return the
+entry at the index of the xa_state.  Using xas_next_entry() to
+move to the next present entry instead of xas_find() will save
+a function call in the majority of cases at the expense of emitting more
+inline code.
+
+The xas_find_marked() function is similar.  If the xa_state has
+not been walked, it will return the entry at the index of the xa_state,
+if it is marked.  Otherwise, it will return the first marked entry after
+the entry referenced by the xa_state.  The xas_next_marked()
+function is the equivalent of xas_next_entry().
+
+When iterating over a range of the XArray using xas_for_each()
+or xas_for_each_marked(), it may be necessary to temporarily stop
+the iteration.  The xas_pause() function exists for this purpose.
+After you have done the necessary work and wish to resume, the xa_state
+is in an appropriate state to continue the iteration after the entry
+you last processed.  If you have interrupts disabled while iterating,
+then it is good manners to pause the iteration and reenable interrupts
+every ``XA_CHECK_SCHED`` entries.
+
+The xas_get_mark(), xas_set_mark() and xas_clear_mark() functions require
+the xa_state cursor to have been moved to the appropriate location in the
+XArray; they will do nothing if you have called xas_pause() or xas_set()
+immediately before.
+
+You can call xas_set_update() to have a callback function
+called each time the XArray updates a node.  This is used by the page
+cache workingset code to maintain its list of nodes which contain only
+shadow entries.
+
+Multi-Index Entries
+-------------------
+
+The XArray has the ability to tie multiple indices together so that
+operations on one index affect all indices.  For example, storing into
+any index will change the value of the entry retrieved from any index.
+Setting or clearing a mark on any index will set or clear the mark
+on every index that is tied together.  The current implementation
+only allows tying ranges which are aligned powers of two together;
+eg indices 64-127 may be tied together, but 2-6 may not be.  This may
+save substantial quantities of memory; for example tying 512 entries
+together will save over 4kB.
+
+You can create a multi-index entry by using XA_STATE_ORDER()
+or xas_set_order() followed by a call to xas_store().
+Calling xas_load() with a multi-index xa_state will walk the
+xa_state to the right location in the tree, but the return value is not
+meaningful, potentially being an internal entry or ``NULL`` even when there
+is an entry stored within the range.  Calling xas_find_conflict()
+will return the first entry within the range or ``NULL`` if there are no
+entries in the range.  The xas_for_each_conflict() iterator will
+iterate over every entry which overlaps the specified range.
+
+If xas_load() encounters a multi-index entry, the xa_index
+in the xa_state will not be changed.  When iterating over an XArray
+or calling xas_find(), if the initial index is in the middle
+of a multi-index entry, it will not be altered.  Subsequent calls
+or iterations will move the index to the first index in the range.
+Each entry will only be returned once, no matter how many indices it
+occupies.
+
+Using xas_next() or xas_prev() with a multi-index xa_state is not
+supported.  Using either of these functions on a multi-index entry will
+reveal sibling entries; these should be skipped over by the caller.
+
+Storing ``NULL`` into any index of a multi-index entry will set the
+entry at every index to ``NULL`` and dissolve the tie.  A multi-index
+entry can be split into entries occupying smaller ranges by calling
+xas_split_alloc() without the xa_lock held, followed by taking the lock
+and calling xas_split().
+
+Functions and structures
+========================
+
+.. kernel-doc:: include/linux/xarray.h
+.. kernel-doc:: lib/xarray.c