--- /dev/null
+<HTML>
+<HEAD>
+ <TITLE> Conservative GC Algorithmic Overview </TITLE>
+ <AUTHOR> Hans-J. Boehm, HP Labs (Some of this was written at SGI)</author>
+</HEAD>
+<BODY>
+<H1> <I>This is under construction, and may always be.</i> </h1>
+<H1> Conservative GC Algorithmic Overview </h1>
+<P>
+This is a description of the algorithms and data structures used in our
+conservative garbage collector. I expect the level of detail to increase
+with time. For a survey of GC algorithms, see for example
+<A HREF="ftp://ftp.cs.utexas.edu/pub/garbage/gcsurvey.ps"> Paul Wilson's
+excellent paper</a>. For an overview of the collector interface,
+see <A HREF="gcinterface.html">here</a>.
+<P>
+This description is targeted primarily at someone trying to understand the
+source code. It specifically refers to variable and function names.
+It may also be useful for understanding the algorithms at a higher level.
+<P>
+The description here assumes that the collector is used in default mode.
+In particular, we assume that it used as a garbage collector, and not just
+a leak detector. We initially assume that it is used in stop-the-world,
+non-incremental mode, though the presence of the incremental collector
+will be apparent in the design.
+We assume the default finalization model, but the code affected by that
+is very localized.
+<H2> Introduction </h2>
+The garbage collector uses a modified mark-sweep algorithm. Conceptually
+it operates roughly in four phases, which are performed occasionally
+as part of a memory allocation:
+
+<OL>
+
+<LI>
+<I>Preparation</i> Each object has an associated mark bit.
+Clear all mark bits, indicating that all objects
+are potentially unreachable.
+
+<LI>
+<I>Mark phase</i> Marks all objects that can be reachable via chains of
+pointers from variables. Often the collector has no real information
+about the location of pointer variables in the heap, so it
+views all static data areas, stacks and registers as potentially containing
+pointers. Any bit patterns that represent addresses inside
+heap objects managed by the collector are viewed as pointers.
+Unless the client program has made heap object layout information
+available to the collector, any heap objects found to be reachable from
+variables are again scanned similarly.
+
+<LI>
+<I>Sweep phase</i> Scans the heap for inaccessible, and hence unmarked,
+objects, and returns them to an appropriate free list for reuse. This is
+not really a separate phase; even in non incremental mode this is operation
+is usually performed on demand during an allocation that discovers an empty
+free list. Thus the sweep phase is very unlikely to touch a page that
+would not have been touched shortly thereafter anyway.
+
+<LI>
+<I>Finalization phase</i> Unreachable objects which had been registered
+for finalization are enqueued for finalization outside the collector.
+
+</ol>
+
+<P>
+The remaining sections describe the memory allocation data structures,
+and then the last 3 collection phases in more detail. We conclude by
+outlining some of the additional features implemented in the collector.
+
+<H2>Allocation</h2>
+The collector includes its own memory allocator. The allocator obtains
+memory from the system in a platform-dependent way. Under UNIX, it
+uses either <TT>malloc</tt>, <TT>sbrk</tt>, or <TT>mmap</tt>.
+<P>
+Most static data used by the allocator, as well as that needed by the
+rest of the garbage collector is stored inside the
+<TT>_GC_arrays</tt> structure.
+This allows the garbage collector to easily ignore the collectors own
+data structures when it searches for root pointers. Other allocator
+and collector internal data structures are allocated dynamically
+with <TT>GC_scratch_alloc</tt>. <TT>GC_scratch_alloc</tt> does not
+allow for deallocation, and is therefore used only for permanent data
+structures.
+<P>
+The allocator allocates objects of different <I>kinds</i>.
+Different kinds are handled somewhat differently by certain parts
+of the garbage collector. Certain kinds are scanned for pointers,
+others are not. Some may have per-object type descriptors that
+determine pointer locations. Or a specific kind may correspond
+to one specific object layout. Two built-in kinds are uncollectable.
+One (<TT>STUBBORN</tt>) is immutable without special precautions.
+In spite of that, it is very likely that most C clients of the
+collector currently
+use at most two kinds: <TT>NORMAL</tt> and <TT>PTRFREE</tt> objects.
+The <A HREF="http://gcc.gnu.org/java">gcj</a> runtime also makes
+heavy use of a kind (allocated with GC_gcj_malloc) that stores
+type information at a known offset in method tables.
+<P>
+The collector uses a two level allocator. A large block is defined to
+be one larger than half of <TT>HBLKSIZE</tt>, which is a power of 2,
+typically on the order of the page size.
+<P>
+Large block sizes are rounded up to
+the next multiple of <TT>HBLKSIZE</tt> and then allocated by
+<TT>GC_allochblk</tt>. Recent versions of the collector
+use an approximate best fit algorithm by keeping free lists for
+several large block sizes.
+The actual
+implementation of <TT>GC_allochblk</tt>
+is significantly complicated by black-listing issues
+(see below).
+<P>
+Small blocks are allocated in chunks of size <TT>HBLKSIZE</tt>.
+Each chunk is
+dedicated to only one object size and kind.
+<P>
+The allocator maintains
+separate free lists for each size and kind of object.
+Associated with each kind is an array of free list pointers,
+with entry <TT>freelist[</tt><I>i</i><TT>]</tt> pointing to
+a free list of size <I>i</i> objects.
+In recent versions of the
+collector, index <TT>i</tt> is expressed in granules, which are the
+minimum allocatable unit, typically 8 or 16 bytes.
+The free lists themselves are
+linked through the first word in each object (see <TT>obj_link()</tt>
+macro).
+<P>
+Once a large block is split for use in smaller objects, it can only
+be used for objects of that size, unless the collector discovers a completely
+empty chunk. Completely empty chunks are restored to the appropriate
+large block free list.
+<P>
+In order to avoid allocating blocks for too many distinct object sizes,
+the collector normally does not directly allocate objects of every possible
+request size. Instead request are rounded up to one of a smaller number
+of allocated sizes, for which free lists are maintained. The exact
+allocated sizes are computed on demand, but subject to the constraint
+that they increase roughly in geometric progression. Thus objects
+requested early in the execution are likely to be allocated with exactly
+the requested size, subject to alignment constraints.
+See <TT>GC_init_size_map</tt> for details.
+<P>
+The actual size rounding operation during small object allocation is
+implemented as a table lookup in <TT>GC_size_map</tt> which maps
+a requested allocation size in bytes to a number of granules.
+<P>
+Both collector initialization and computation of allocated sizes are
+handled carefully so that they do not slow down the small object fast
+allocation path. An attempt to allocate before the collector is initialized,
+or before the appropriate <TT>GC_size_map</tt> entry is computed,
+will take the same path as an allocation attempt with an empty free list.
+This results in a call to the slow path code (<TT>GC_generic_malloc_inner</tt>)
+which performs the appropriate initialization checks.
+<P>
+In non-incremental mode, we make a decision about whether to garbage collect
+whenever an allocation would otherwise have failed with the current heap size.
+If the total amount of allocation since the last collection is less than
+the heap size divided by <TT>GC_free_space_divisor</tt>, we try to
+expand the heap. Otherwise, we initiate a garbage collection. This ensures
+that the amount of garbage collection work per allocated byte remains
+constant.
+<P>
+The above is in fact an oversimplification of the real heap expansion
+and GC triggering heuristic, which adjusts slightly for root size
+and certain kinds of
+fragmentation. In particular:
+<UL>
+<LI> Programs with a large root set size and
+little live heap memory will expand the heap to amortize the cost of
+scanning the roots.
+<LI> Versions 5.x of the collector actually collect more frequently in
+nonincremental mode. The large block allocator usually refuses to split
+large heap blocks once the garbage collection threshold is
+reached. This often has the effect of collecting well before the
+heap fills up, thus reducing fragmentation and working set size at the
+expense of GC time. Versions 6.x choose an intermediate strategy depending
+on how much large object allocation has taken place in the past.
+(If the collector is configured to unmap unused pages, versions 6.x
+use the 5.x strategy.)
+<LI> In calculating the amount of allocation since the last collection we
+give partial credit for objects we expect to be explicitly deallocated.
+Even if all objects are explicitly managed, it is often desirable to collect
+on rare occasion, since that is our only mechanism for coalescing completely
+empty chunks.
+</ul>
+<P>
+It has been suggested that this should be adjusted so that we favor
+expansion if the resulting heap still fits into physical memory.
+In many cases, that would no doubt help. But it is tricky to do this
+in a way that remains robust if multiple application are contending
+for a single pool of physical memory.
+
+<H2>Mark phase</h2>
+
+At each collection, the collector marks all objects that are
+possibly reachable from pointer variables. Since it cannot generally
+tell where pointer variables are located, it scans the following
+<I>root segments</i> for pointers:
+<UL>
+<LI>The registers. Depending on the architecture, this may be done using
+assembly code, or by calling a <TT>setjmp</tt>-like function which saves
+register contents on the stack.
+<LI>The stack(s). In the case of a single-threaded application,
+on most platforms this
+is done by scanning the memory between (an approximation of) the current
+stack pointer and <TT>GC_stackbottom</tt>. (For Itanium, the register stack
+scanned separately.) The <TT>GC_stackbottom</tt> variable is set in
+a highly platform-specific way depending on the appropriate configuration
+information in <TT>gcconfig.h</tt>. Note that the currently active
+stack needs to be scanned carefully, since callee-save registers of
+client code may appear inside collector stack frames, which may
+change during the mark process. This is addressed by scanning
+some sections of the stack "eagerly", effectively capturing a snapshot
+at one point in time.
+<LI>Static data region(s). In the simplest case, this is the region
+between <TT>DATASTART</tt> and <TT>DATAEND</tt>, as defined in
+<TT>gcconfig.h</tt>. However, in most cases, this will also involve
+static data regions associated with dynamic libraries. These are
+identified by the mostly platform-specific code in <TT>dyn_load.c</tt>.
+</ul>
+The marker maintains an explicit stack of memory regions that are known
+to be accessible, but that have not yet been searched for contained pointers.
+Each stack entry contains the starting address of the block to be scanned,
+as well as a descriptor of the block. If no layout information is
+available for the block, then the descriptor is simply a length.
+(For other possibilities, see <TT>gc_mark.h</tt>.)
+<P>
+At the beginning of the mark phase, all root segments
+(as described above) are pushed on the
+stack by <TT>GC_push_roots</tt>. (Registers and eagerly processed
+stack sections are processed by pushing the referenced objects instead
+of the stack section itself.) If <TT>ALL_INTERIOR_PTRS</tt> is not
+defined, then stack roots require special treatment. In this case, the
+normal marking code ignores interior pointers, but <TT>GC_push_all_stack</tt>
+explicitly checks for interior pointers and pushes descriptors for target
+objects.
+<P>
+The marker is structured to allow incremental marking.
+Each call to <TT>GC_mark_some</tt> performs a small amount of
+work towards marking the heap.
+It maintains
+explicit state in the form of <TT>GC_mark_state</tt>, which
+identifies a particular sub-phase. Some other pieces of state, most
+notably the mark stack, identify how much work remains to be done
+in each sub-phase. The normal progression of mark states for
+a stop-the-world collection is:
+<OL>
+<LI> <TT>MS_INVALID</tt> indicating that there may be accessible unmarked
+objects. In this case <TT>GC_objects_are_marked</tt> will simultaneously
+be false, so the mark state is advanced to
+<LI> <TT>MS_PUSH_UNCOLLECTABLE</tt> indicating that it suffices to push
+uncollectable objects, roots, and then mark everything reachable from them.
+<TT>Scan_ptr</tt> is advanced through the heap until all uncollectable
+objects are pushed, and objects reachable from them are marked.
+At that point, the next call to <TT>GC_mark_some</tt> calls
+<TT>GC_push_roots</tt> to push the roots. It the advances the
+mark state to
+<LI> <TT>MS_ROOTS_PUSHED</tt> asserting that once the mark stack is
+empty, all reachable objects are marked. Once in this state, we work
+only on emptying the mark stack. Once this is completed, the state
+changes to
+<LI> <TT>MS_NONE</tt> indicating that reachable objects are marked.
+</ol>
+
+The core mark routine <TT>GC_mark_from</tt>, is called
+repeatedly by several of the sub-phases when the mark stack starts to fill
+up. It is also called repeatedly in <TT>MS_ROOTS_PUSHED</tt> state
+to empty the mark stack.
+The routine is designed to only perform a limited amount of marking at
+each call, so that it can also be used by the incremental collector.
+It is fairly carefully tuned, since it usually consumes a large majority
+of the garbage collection time.
+<P>
+The fact that it perform a only a small amount of work per call also
+allows it to be used as the core routine of the parallel marker. In that
+case it is normally invoked on thread-private mark stacks instead of the
+global mark stack. More details can be found in
+<A HREF="scale.html">scale.html</a>
+<P>
+The marker correctly handles mark stack overflows. Whenever the mark stack
+overflows, the mark state is reset to <TT>MS_INVALID</tt>.
+Since there are already marked objects in the heap,
+this eventually forces a complete
+scan of the heap, searching for pointers, during which any unmarked objects
+referenced by marked objects are again pushed on the mark stack. This
+process is repeated until the mark phase completes without a stack overflow.
+Each time the stack overflows, an attempt is made to grow the mark stack.
+All pieces of the collector that push regions onto the mark stack have to be
+careful to ensure forward progress, even in case of repeated mark stack
+overflows. Every mark attempt results in additional marked objects.
+<P>
+Each mark stack entry is processed by examining all candidate pointers
+in the range described by the entry. If the region has no associated
+type information, then this typically requires that each 4-byte aligned
+quantity (8-byte aligned with 64-bit pointers) be considered a candidate
+pointer.
+<P>
+We determine whether a candidate pointer is actually the address of
+a heap block. This is done in the following steps:
+<NL>
+<LI> The candidate pointer is checked against rough heap bounds.
+These heap bounds are maintained such that all actual heap objects
+fall between them. In order to facilitate black-listing (see below)
+we also include address regions that the heap is likely to expand into.
+Most non-pointers fail this initial test.
+<LI> The candidate pointer is divided into two pieces; the most significant
+bits identify a <TT>HBLKSIZE</tt>-sized page in the address space, and
+the least significant bits specify an offset within that page.
+(A hardware page may actually consist of multiple such pages.
+HBLKSIZE is usually the page size divided by a small power of two.)
+<LI>
+The page address part of the candidate pointer is looked up in a
+<A HREF="tree.html">table</a>.
+Each table entry contains either 0, indicating that the page is not part
+of the garbage collected heap, a small integer <I>n</i>, indicating
+that the page is part of large object, starting at least <I>n</i> pages
+back, or a pointer to a descriptor for the page. In the first case,
+the candidate pointer i not a true pointer and can be safely ignored.
+In the last two cases, we can obtain a descriptor for the page containing
+the beginning of the object.
+<LI>
+The starting address of the referenced object is computed.
+The page descriptor contains the size of the object(s)
+in that page, the object kind, and the necessary mark bits for those
+objects. The size information can be used to map the candidate pointer
+to the object starting address. To accelerate this process, the page header
+also contains a pointer to a precomputed map of page offsets to displacements
+from the beginning of an object. The use of this map avoids a
+potentially slow integer remainder operation in computing the object
+start address.
+<LI>
+The mark bit for the target object is checked and set. If the object
+was previously unmarked, the object is pushed on the mark stack.
+The descriptor is read from the page descriptor. (This is computed
+from information <TT>GC_obj_kinds</tt> when the page is first allocated.)
+</nl>
+<P>
+At the end of the mark phase, mark bits for left-over free lists are cleared,
+in case a free list was accidentally marked due to a stray pointer.
+
+<H2>Sweep phase</h2>
+
+At the end of the mark phase, all blocks in the heap are examined.
+Unmarked large objects are immediately returned to the large object free list.
+Each small object page is checked to see if all mark bits are clear.
+If so, the entire page is returned to the large object free list.
+Small object pages containing some reachable object are queued for later
+sweeping, unless we determine that the page contains very little free
+space, in which case it is not examined further.
+<P>
+This initial sweep pass touches only block headers, not
+the blocks themselves. Thus it does not require significant paging, even
+if large sections of the heap are not in physical memory.
+<P>
+Nonempty small object pages are swept when an allocation attempt
+encounters an empty free list for that object size and kind.
+Pages for the correct size and kind are repeatedly swept until at
+least one empty block is found. Sweeping such a page involves
+scanning the mark bit array in the page header, and building a free
+list linked through the first words in the objects themselves.
+This does involve touching the appropriate data page, but in most cases
+it will be touched only just before it is used for allocation.
+Hence any paging is essentially unavoidable.
+<P>
+Except in the case of pointer-free objects, we maintain the invariant
+that any object in a small object free list is cleared (except possibly
+for the link field). Thus it becomes the burden of the small object
+sweep routine to clear objects. This has the advantage that we can
+easily recover from accidentally marking a free list, though that could
+also be handled by other means. The collector currently spends a fair
+amount of time clearing objects, and this approach should probably be
+revisited.
+<P>
+In most configurations, we use specialized sweep routines to handle common
+small object sizes. Since we allocate one mark bit per word, it becomes
+easier to examine the relevant mark bits if the object size divides
+the word length evenly. We also suitably unroll the inner sweep loop
+in each case. (It is conceivable that profile-based procedure cloning
+in the compiler could make this unnecessary and counterproductive. I
+know of no existing compiler to which this applies.)
+<P>
+The sweeping of small object pages could be avoided completely at the expense
+of examining mark bits directly in the allocator. This would probably
+be more expensive, since each allocation call would have to reload
+a large amount of state (e.g. next object address to be swept, position
+in mark bit table) before it could do its work. The current scheme
+keeps the allocator simple and allows useful optimizations in the sweeper.
+
+<H2>Finalization</h2>
+Both <TT>GC_register_disappearing_link</tt> and
+<TT>GC_register_finalizer</tt> add the request to a corresponding hash
+table. The hash table is allocated out of collected memory, but
+the reference to the finalizable object is hidden from the collector.
+Currently finalization requests are processed non-incrementally at the
+end of a mark cycle.
+<P>
+The collector makes an initial pass over the table of finalizable objects,
+pushing the contents of unmarked objects onto the mark stack.
+After pushing each object, the marker is invoked to mark all objects
+reachable from it. The object itself is not explicitly marked.
+This assures that objects on which a finalizer depends are neither
+collected nor finalized.
+<P>
+If in the process of marking from an object the
+object itself becomes marked, we have uncovered
+a cycle involving the object. This usually results in a warning from the
+collector. Such objects are not finalized, since it may be
+unsafe to do so. See the more detailed
+<A HREF="http://www.hpl.hp.com/personal/Hans_Boehm/gc/finalization.html"> discussion of finalization semantics</a>.
+<P>
+Any objects remaining unmarked at the end of this process are added to
+a queue of objects whose finalizers can be run. Depending on collector
+configuration, finalizers are dequeued and run either implicitly during
+allocation calls, or explicitly in response to a user request.
+(Note that the former is unfortunately both the default and not generally safe.
+If finalizers perform synchronization, it may result in deadlocks.
+Nontrivial finalizers generally need to perform synchronization, and
+thus require a different collector configuration.)
+<P>
+The collector provides a mechanism for replacing the procedure that is
+used to mark through objects. This is used both to provide support for
+Java-style unordered finalization, and to ignore certain kinds of cycles,
+<I>e.g.</i> those arising from C++ implementations of virtual inheritance.
+
+<H2>Generational Collection and Dirty Bits</h2>
+We basically use the concurrent and generational GC algorithm described in
+<A HREF="http://www.hpl.hp.com/personal/Hans_Boehm/gc/papers/pldi91.ps.Z">"Mostly Parallel Garbage Collection"</a>,
+by Boehm, Demers, and Shenker.
+<P>
+The most significant modification is that
+the collector always starts running in the allocating thread.
+There is no separate garbage collector thread. (If parallel GC is
+enabled, helper threads may also be woken up.)
+If an allocation attempt either requests a large object, or encounters
+an empty small object free list, and notices that there is a collection
+in progress, it immediately performs a small amount of marking work
+as described above.
+<P>
+This change was made both because we wanted to easily accommodate
+single-threaded environments, and because a separate GC thread requires
+very careful control over the scheduler to prevent the mutator from
+out-running the collector, and hence provoking unneeded heap growth.
+<P>
+In incremental mode, the heap is always expanded when we encounter
+insufficient space for an allocation. Garbage collection is triggered
+whenever we notice that more than
+<TT>GC_heap_size</tt>/2 * <TT>GC_free_space_divisor</tt>
+bytes of allocation have taken place.
+After <TT>GC_full_freq</tt> minor collections a major collection
+is started.
+<P>
+All collections initially run uninterrupted until a predetermined
+amount of time (50 msecs by default) has expired. If this allows
+the collection to complete entirely, we can avoid correcting
+for data structure modifications during the collection. If it does
+not complete, we return control to the mutator, and perform small
+amounts of additional GC work during those later allocations that
+cannot be satisfied from small object free lists. When marking completes,
+the set of modified pages is retrieved, and we mark once again from
+marked objects on those pages, this time with the mutator stopped.
+<P>
+We keep track of modified pages using one of several distinct mechanisms:
+<OL>
+<LI>
+Through explicit mutator cooperation. Currently this requires
+the use of <TT>GC_malloc_stubborn</tt>, and is rarely used.
+<LI>
+(<TT>MPROTECT_VDB</tt>) By write-protecting physical pages and
+catching write faults. This is
+implemented for many Unix-like systems and for win32. It is not possible
+in a few environments.
+<LI>
+(<TT>PROC_VDB</tt>) By retrieving dirty bit information from /proc.
+(Currently only Sun's
+Solaris supports this. Though this is considerably cleaner, performance
+may actually be better with mprotect and signals.)
+<LI>
+(<TT>PCR_VDB</tt>) By relying on an external dirty bit implementation, in this
+case the one in Xerox PCR.
+<LI>
+(<TT>DEFAULT_VDB</tt>) By treating all pages as dirty. This is the default if
+none of the other techniques is known to be usable, and
+<TT>GC_malloc_stubborn</tt> is not used. Practical only for testing, or if
+the vast majority of objects use <TT>GC_malloc_stubborn</tt>.
+</ol>
+
+<H2>Black-listing</h2>
+
+The collector implements <I>black-listing</i> of pages, as described
+in
+<A HREF="http://www.acm.org/pubs/citations/proceedings/pldi/155090/p197-boehm/">
+Boehm, ``Space Efficient Conservative Collection'', PLDI '93</a>, also available
+<A HREF="papers/pldi93.ps.Z">here</a>.
+<P>
+During the mark phase, the collector tracks ``near misses'', i.e. attempts
+to follow a ``pointer'' to just outside the garbage-collected heap, or
+to a currently unallocated page inside the heap. Pages that have been
+the targets of such near misses are likely to be the targets of
+misidentified ``pointers'' in the future. To minimize the future
+damage caused by such misidentifications they will be allocated only to
+small pointerfree objects.
+<P>
+The collector understands two different kinds of black-listing. A
+page may be black listed for interior pointer references
+(<TT>GC_add_to_black_list_stack</tt>), if it was the target of a near
+miss from a location that requires interior pointer recognition,
+<I>e.g.</i> the stack, or the heap if <TT>GC_all_interior_pointers</tt>
+is set. In this case, we also avoid allocating large blocks that include
+this page.
+<P>
+If the near miss came from a source that did not require interior
+pointer recognition, it is black-listed with
+<TT>GC_add_to_black_list_normal</tt>.
+A page black-listed in this way may appear inside a large object,
+so long as it is not the first page of a large object.
+<P>
+The <TT>GC_allochblk</tt> routine respects black-listing when assigning
+a block to a particular object kind and size. It occasionally
+drops (i.e. allocates and forgets) blocks that are completely black-listed
+in order to avoid excessively long large block free lists containing
+only unusable blocks. This would otherwise become an issue
+if there is low demand for small pointerfree objects.
+
+<H2>Thread support</h2>
+We support several different threading models. Unfortunately Pthreads,
+the only reasonably well standardized thread model, supports too narrow
+an interface for conservative garbage collection. There appears to be
+no completely portable way to allow the collector
+to coexist with various Pthreads
+implementations. Hence we currently support only the more
+common Pthreads implementations.
+<P>
+In particular, it is very difficult for the collector to stop all other
+threads in the system and examine the register contents. This is currently
+accomplished with very different mechanisms for some Pthreads
+implementations. The Solaris implementation temporarily disables much
+of the user-level threads implementation by stopping kernel-level threads
+("lwp"s). The Linux/HPUX/OSF1 and Irix implementations sends signals to
+individual Pthreads and has them wait in the signal handler.
+<P>
+The Linux and Irix implementations use
+only documented Pthreads calls, but rely on extensions to their semantics.
+The Linux implementation <TT>linux_threads.c</tt> relies on only very
+mild extensions to the pthreads semantics, and already supports a large number
+of other Unix-like pthreads implementations. Our goal is to make this the
+only pthread support in the collector.
+<P>
+(The Irix implementation is separate only for historical reasons and should
+clearly be merged. The current Solaris implementation probably performs
+better in the uniprocessor case, but does not support thread operations in the
+collector. Hence it cannot support the parallel marker.)
+<P>
+All implementations must
+intercept thread creation and a few other thread-specific calls to allow
+enumeration of threads and location of thread stacks. This is current
+accomplished with <TT># define</tt>'s in <TT>gc.h</tt>
+(really <TT>gc_pthread_redirects.h</tt>), or optionally
+by using ld's function call wrapping mechanism under Linux.
+<P>
+Recent versions of the collector support several facilities to enhance
+the processor-scalability and thread performance of the collector.
+These are discussed in more detail <A HREF="scale.html">here</a>.
+We briefly outline the data approach to thread-local allocation in the
+next section.
+<H2>Thread-local allocation</h2>
+If thread-local allocation is enabled, the collector keeps separate
+arrays of free lists for each thread. Thread-local allocation
+is currently only supported on a few platforms.
+<P>
+The free list arrays associated
+with each thread are only used to satisfy requests for objects that
+are both very small, and belong to one of a small number of well-known
+kinds. These currently include "normal" and pointer-free objects.
+Depending on the configuration, "gcj" objects may also be included.
+<P>
+Thread-local free list entries contain either a pointer to the first
+element of a free list, or they contain a counter of the number of
+allocation granules, corresponding to objects of this size,
+allocated so far. Initially they contain the
+value one, i.e. a small counter value.
+<P>
+Thread-local allocation allocates directly through the global
+allocator, if the object is of a size or kind not covered by the
+local free lists.
+<P>
+If there is an appropriate local free list, the allocator checks whether it
+contains a sufficiently small counter value. If so, the counter is simply
+incremented by the counter value, and the global allocator is used.
+In this way, the initial few allocations of a given size bypass the local
+allocator. A thread that only allocates a handful of objects of a given
+size will not build up its own free list for that size. This avoids
+wasting space for unpopular objects sizes or kinds.
+<P>
+Once the counter passes a threshold, <TT>GC_malloc_many</tt> is called
+to allocate roughly <TT>HBLKSIZE</tt> space and put it on the corresponding
+local free list. Further allocations of that size and kind then use
+this free list, and no longer need to acquire the allocation lock.
+The allocation procedure is otherwise similar to the global free lists.
+The local free lists are also linked using the first word in the object.
+In most cases this means they require considerably less time.
+<P>
+Local free lists are treated buy most of the rest of the collector
+as though they were in-use reachable data. This requires some care,
+since pointer-free objects are not normally traced, and hence a special
+tracing procedure is required to mark all objects on pointer-free and
+gcj local free lists.
+<P>
+On thread exit, any remaining thread-local free list entries are
+transferred back to the global free list.
+<P>
+Note that if the collector is configured for thread-local allocation,
+GC versions before 7 do not invoke the thread-local allocator by default.
+<TT>GC_malloc</tt> only uses thread-local allocation in version 7 and later.
+<P>
+For some more details see <A HREF="scale.html">here</a>, and the
+technical report entitled
+<A HREF="http://www.hpl.hp.com/techreports/2000/HPL-2000-165.html">
+``Fast Multiprocessor Memory Allocation and Garbage Collection''
+</a>
+<P>
+<HR>
+<P>
+Comments are appreciated. Please send mail to
+<A HREF="mailto:gc@linux.hpl.hp.com"><tt>gc@linux.hpl.hp.com</tt></a>
+(GC mailing list) or
+<A HREF="mailto:Hans.Boehm@hp.com"><TT>Hans.Boehm@hp.com</tt></a>
+<P>
+This is a modified copy of a page written while the author was at SGI.
+The original was <A HREF="http://reality.sgi.com/boehm/gcdescr.html">here</a>.
+</body>
+</html>
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