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+ <TITLE> Conservative GC Algorithmic Overview </TITLE>
+ <AUTHOR> Hans-J. Boehm, Silicon Graphics</author>
+<H1> <I>This is under construction</i> </h1>
+<H1> Conservative GC Algorithmic Overview </h1>
+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>.
+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.
+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:
+<I>Preparation</i> Clear all mark bits, indicating that all objects
+are potentially unreachable.
+<I>Mark phase</i> Marks all objects that can be reachable via chains of
+pointers from variables. Normally 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
+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.
+<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.
+<I>Finalization phase</i> Unreachable objects which had been registered
+for finalization are enqueued for finalization outside the collector.
+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.
+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>.
+Most static data used by the allocator, as well as that needed by the
+rest of the garbage collector is stored inside the
+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
+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 applications currently
+use at most two kinds: <TT>NORMAL</tt> and <TT>PTRFREE</tt> objects.
+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.
+Large block sizes are rounded up to
+the next multiple of <TT>HBLKSIZE</tt> and then allocated by
+<TT>GC_allochblk</tt>. This uses roughly what Paul Wilson has termed
+a "next fit" algorithm, i.e. first-fit with a rotating pointer.
+The implementation does check for a better fitting immediately
+adjacent block, which gives it somewhat better fragmentation characteristics.
+I'm now convinced it should use a best fit algorithm. The actual
+implementation of <TT>GC_allochblk</tt>
+is significantly complicated by black-listing issues
+Small blocks are allocated in blocks of size <TT>HBLKSIZE</tt>.
+Each block is
+dedicated to only one object size and kind. The allocator maintains
+separate free lists for each size and kind of object.
+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.
+The actual size rounding operation during small object allocation is
+implemented as a table lookup in <TT>GC_size_map</tt>.
+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.
+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
+The above is in fat an oversimplification of the real heap expansion
+heuristic, which adjusts slightly for root size and certain kinds of
+fragmentation. In particular, programs with a large root set size and
+little live heap memory will expand the heap to amortize the cost of
+scanning the roots.
+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. 6.x will chose 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
+will use the 5.x strategy.)
+(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.)
+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>.)
+At the beginning of the mark phase, all root segments are pushed on the
+stack by <TT>GC_push_roots</tt>. 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
+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.
+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:
+<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
+<LI> <TT>MS_NONE</tt> indicating that reachable objects are marked.
+The core mark routine <TT>GC_mark_from_mark_stack</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.
+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.
+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
+We determine whether a candidate pointer is actually the address of
+a heap block. This is done in the following steps:
+<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.)
+The page address part of the candidate pointer is looked up in 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.
+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
+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.)
+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.
+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
+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.
+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.
+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
+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.)
+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.
+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.
+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.
+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="finalization.html"> discussion of finalization semantics</a>.
+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.
+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 parallel and generational GC algorithm described in
+<A HREF="papers/pldi91.ps.gz">"Mostly Parallel Garbage Collection"</a>,
+by Boehm, Demers, and Shenker.
+The most significant modification is that
+the collector always runs in the allocating thread.
+There is no separate garbage collector thread.
+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.
+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.
+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
+All collections initially run interrupted 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.
+We keep track of modified pages using one of three distinct mechanisms:
+Through explicit mutator cooperation. Currently this requires
+the use of <TT>GC_malloc_stubborn</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.
+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.)
+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 portable way to allow the collector to coexist with various Pthreads
+implementations. Hence we currently support only a few of the more
+common Pthreads implementations.
+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 different Pthreads
+implementations. The Solaris implementation temporarily disables much
+of the user-level threads implementation by stopping kernel-level threads
+("lwp"s). The Irix implementation sends signals to individual Pthreads
+and has them wait in the signal handler. The Linux implementation
+is similar in spirit to the Irix one.
+The Irix implementation uses
+only documented Pthreads calls, but relies on extensions to their semantics,
+notably the use of mutexes and condition variables from signal
+handlers. The Linux implementation should be far closer to
+portable, though impirically it is not completely portable.
+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>, or optionally
+by using ld's function call wrapping mechanism under Linux.
+Comments are appreciated. Please send mail to