%% ``The contents of this file are subject to the Erlang Public License,
%% Version 1.1, (the "License"); you may not use this file except in
%% compliance with the License. You should have received a copy of the
%% Erlang Public License along with this software. If not, it can be
%% retrieved via the world wide web at http://www.erlang.org/.
%%
%% Software distributed under the License is distributed on an "AS IS"
%% basis, WITHOUT WARRANTY OF ANY KIND, either express or implied. See
%% the License for the specific language governing rights and limitations
%% under the License.
%%
%% The Initial Developer of the Original Code is Richard Carlsson.
%% Copyright (C) 1999-2002 Richard Carlsson.
%% Portions created by Ericsson are Copyright 2001, Ericsson Utvecklings
%% AB. All Rights Reserved.''
%%
%% $Id: cerl_inline.erl,v 1.1 2008/12/17 09:53:41 mikpe Exp $
%%
%% Core Erlang inliner.
%% =====================================================================
%%
%% This is an implementation of the algorithm by Waddell and Dybvig
%% ("Fast and Effective Procedure Inlining", International Static
%% Analysis Symposium 1997), adapted to the Core Erlang language.
%%
%% Instead of always renaming variables and function variables, this
%% implementation uses the "no-shadowing strategy" of Peyton Jones and
%% Marlow ("Secrets of the Glasgow Haskell Compiler Inliner", 1999).
%%
%% =====================================================================
%% TODO: inline single-source-reference operands without size limit.
-module(cerl_inline).
-export([core_transform/2, transform/1, transform/2]).
-import(cerl, [abstract/1, alias_pat/1, alias_var/1, apply_args/1,
apply_op/1, atom_name/1, atom_val/1, bitstr_val/1,
bitstr_size/1, bitstr_unit/1, bitstr_type/1,
bitstr_flags/1, binary_segments/1, update_c_alias/3,
update_c_apply/3, update_c_binary/2, update_c_bitstr/6,
update_c_call/4, update_c_case/3, update_c_catch/2,
update_c_clause/4, c_fun/2, c_int/1, c_let/3,
update_c_let/4, update_c_letrec/3, update_c_module/5,
update_c_primop/3, update_c_receive/4, update_c_seq/3,
c_seq/2, update_c_try/6, c_tuple/1, update_c_values/2,
c_values/1, c_var/1, call_args/1, call_module/1,
call_name/1, case_arity/1, case_arg/1, case_clauses/1,
catch_body/1, clause_body/1, clause_guard/1,
clause_pats/1, clause_vars/1, concrete/1, cons_hd/1,
cons_tl/1, data_arity/1, data_es/1, data_type/1,
fun_body/1, fun_vars/1, get_ann/1, int_val/1,
is_c_atom/1, is_c_cons/1, is_c_fun/1, is_c_int/1,
is_c_list/1, is_c_seq/1, is_c_tuple/1, is_c_var/1,
is_data/1, is_literal/1, is_literal_term/1, let_arg/1,
let_body/1, let_vars/1, letrec_body/1, letrec_defs/1,
list_length/1, list_elements/1, update_data/3,
make_list/1, make_data_skel/2, module_attrs/1,
module_defs/1, module_exports/1, module_name/1,
primop_args/1, primop_name/1, receive_action/1,
receive_clauses/1, receive_timeout/1, seq_arg/1,
seq_body/1, set_ann/2, try_arg/1, try_body/1, try_vars/1,
try_evars/1, try_handler/1, tuple_es/1, tuple_arity/1,
type/1, values_es/1, var_name/1]).
-import(lists, [foldl/3, foldr/3, mapfoldl/3, reverse/1]).
%%
%% Constants
%%
debug_runtime() -> false.
debug_counters() -> false.
%% Normal execution times for inlining are between 0.1 and 0.3 seconds
%% (on the author's current equipment). The default effort limit of 150
%% is high enough that most normal programs never hit the limit even
%% once, and for difficult programs, it generally keeps the execution
%% times below 2-5 seconds. Using an effort counter of 1000 will thus
%% have no further effect on most programs, but some programs may take
%% as much as 10 seconds or more. Effort counts larger than 2500 have
%% never been observed even on very ill-conditioned programs.
%%
%% Size limits between 6 and 18 tend to actually shrink the code,
%% because of the simplifications made possible by inlining. A limit of
%% 16 seems to be optimal for this purpose, often shrinking the
%% executable code by up to 10%. Size limits between 18 and 30 generally
%% give the same code size as if no inlining was done (i.e., code
%% duplication balances out the simplifications at these levels). A size
%% limit between 1 and 5 tends to inline small functions and propagate
%% constants, but does not cause much simplifications do be done, so the
%% net effect will be a slight increase in code size. For size limits
%% above 30, the executable code size tends to increase with about 10%
%% per 100 units, with some variations depending on the sizes of
%% functions in the source code.
%%
%% Typically, about 90% of the maximum speedup achievable is already
%% reached using a size limit of 30, and 98% is reached at limits around
%% 100-150; there is rarely any point in letting the code size increase
%% by more than 10-15%. If too large functions are inlined, cache
%% effects will slow the program down.
default_effort() -> 150.
default_size() -> 24.
%% Base costs/weights for different kinds of expressions. If these are
%% modified, the size limits above may have to be adjusted.
weight(var) -> 0; % We count no cost for variable accesses.
weight(values) -> 0; % Value aggregates have no cost in themselves.
weight(literal) -> 1; % We assume efficient handling of constants.
weight(data) -> 1; % Base cost; add 1 per element.
weight(element) -> 1; % Cost of storing/fetching an element.
weight(argument) -> 1; % Cost of passing a function argument.
weight('fun') -> 6; % Base cost + average number of free vars.
weight('let') -> 0; % Count no cost for let-bindings.
weight(letrec) -> 0; % Like a let-binding.
weight('case') -> 0; % Case switches have no base cost.
weight(clause) -> 1; % Count one jump at the end of each clause body.
weight('receive') -> 9; % Initialization/cleanup cost.
weight('try') -> 1; % Assume efficient implementation.
weight('catch') -> 1; % See `try'.
weight(apply) -> 3; % Average base cost: call/return.
weight(call) -> 3; % Assume remote-calls as efficient as `apply'.
weight(primop) -> 2; % Assume more efficient than `apply'.
weight(binary) -> 4; % Initialisation base cost.
weight(bitstr) -> 3; % Coding/decoding a value; like a primop.
weight(module) -> 1. % Like a letrec with a constant body
%% These "reference" structures are used for variables and function
%% variables. They keep track of the variable name, any bound operand,
%% and the associated store location.
-record(ref, {name, opnd, loc}).
%% Operand structures contain the operand expression, the renaming and
%% environment, the state location, and the effort counter at the call
%% site (cf. `visit').
-record(opnd, {expr, ren, env, loc, effort}).
%% Since expressions are only visited in `effect' context when they are
%% not bound to a referenced variable, only expressions visited in
%% 'value' context are cached.
-record(cache, {expr, size}).
%% The context flags for an application structure are kept separate from
%% the structure itself. Note that the original algorithm had exactly
%% one operand in each application context structure, while we can have
%% several, or none.
-record(app, {opnds, ctxt, loc}).
%%
%% Interface functions
%%
%% Use compile option `{core_transform, inline}' to insert this as a
%% compilation pass.
core_transform(Code, Opts) ->
cerl:to_records(transform(cerl:from_records(Code), Opts)).
transform(Tree) ->
transform(Tree, []).
transform(Tree, Opts) ->
main(Tree, value, Opts).
main(Tree, Ctxt, Opts) ->
%% We spawn a new process to do the work, so we don't have to worry
%% about cluttering the process dictionary with debugging info, or
%% proper deallocation of ets-tables.
Opts1 = Opts ++ [{inline_size, default_size()},
{inline_effort, default_effort()}],
Reply = self(),
Pid = spawn_link(fun () -> start(Reply, Tree, Ctxt, Opts1) end),
receive
{Pid1, Tree1} when Pid1 == Pid ->
Tree1
end.
start(Reply, Tree, Ctxt, Opts) ->
init_debug(),
case debug_runtime() of
true ->
put(inline_start_time,
element(1, erlang:statistics(runtime)));
_ ->
ok
end,
Size = max(1, proplists:get_value(inline_size, Opts)),
Effort = max(1, proplists:get_value(inline_effort, Opts)),
case proplists:get_bool(verbose, Opts) of
true ->
io:fwrite("Inlining: inline_size=~w inline_effort=~w\n",
[Size, Effort]);
false ->
ok
end,
%% Note that the counters of the new state are passive.
S = st__new(Effort, Size),
%%% Initialization is not needed at present. Note that the code in
%%% `inline_init' is not up-to-date with this module.
%%% {Tree1, S1} = inline_init:init(Tree, S),
%%% {Tree2, _S2} = i(Tree1, Ctxt, S1),
{Tree2, _S2} = i(Tree, Ctxt, S),
report_debug(),
Reply ! {self(), Tree2}.
init_debug() ->
case debug_counters() of
true ->
put(counter_effort_triggers, 0),
put(counter_effort_max, 0),
put(counter_size_triggers, 0),
put(counter_size_max, 0);
_ ->
ok
end.
report_debug() ->
case debug_runtime() of
true ->
{Time, _} = erlang:statistics(runtime),
report("Total run time for inlining: ~.2.0f s.\n",
[(Time - get(inline_start_time))/1000]);
_ ->
ok
end,
case debug_counters() of
true ->
counter_stats();
_ ->
ok
end.
counter_stats() ->
T1 = get(counter_effort_triggers),
T2 = get(counter_size_triggers),
E = get(counter_effort_max),
S = get(counter_size_max),
M1 = io_lib:fwrite("\tNumber of triggered "
"effort counters: ~p.\n", [T1]),
M2 = io_lib:fwrite("\tNumber of triggered "
"size counters: ~p.\n", [T2]),
M3 = io_lib:fwrite("\tLargest active effort counter: ~p.\n",
[E]),
M4 = io_lib:fwrite("\tLargest active size counter: ~p.\n",
[S]),
report("Counter statistics:\n~s", [[M1, M2, M3, M4]]).
%% =====================================================================
%% The main inlining function
%%
%% i(E :: coreErlang(),
%% Ctxt :: value | effect | #app{}
%% Ren :: renaming(),
%% Env :: environment(),
%% S :: state())
%% -> {E', S'}
%%
%% Note: It is expected that the input source code ('E') does not
%% contain free variables. If it does, there is a risk of accidental
%% name capture, in case a generated "new" variable name happens to be
%% the same as the name of a variable that is free further below in the
%% tree; the algorithm only consults the current environment to check if
%% a name already exists.
%%
%% The renaming maps names of source-code variable and function
%% variables to new names as necessary to avoid clashes, according to
%% the "no-shadowing" strategy. The environment maps *residual-code*
%% variables and function variables to operands and global information.
%% Separating the renaming from the environment, and using the
%% residual-code variables instead of the source-code variables as its
%% domain, improves the behaviour of the algorithm when code needs to be
%% traversed more than once.
%%
%% Note that there is no such thing as a `test' context for expressions
%% in (Core) Erlang (see `i_case' below for details).
i(E, Ctxt, S) ->
i(E, Ctxt, ren__identity(), env__empty(), S).
i(E, Ctxt, Ren, Env, S0) ->
%% Count one unit of effort on each pass.
S = count_effort(1, S0),
case is_data(E) of
true ->
i_data(E, Ctxt, Ren, Env, S);
false ->
case type(E) of
var ->
i_var(E, Ctxt, Ren, Env, S);
values ->
i_values(E, Ctxt, Ren, Env, S);
'fun' ->
i_fun(E, Ctxt, Ren, Env, S);
seq ->
i_seq(E, Ctxt, Ren, Env, S);
'let' ->
i_let(E, Ctxt, Ren, Env, S);
letrec ->
i_letrec(E, Ctxt, Ren, Env, S);
'case' ->
i_case(E, Ctxt, Ren, Env, S);
'receive' ->
i_receive(E, Ctxt, Ren, Env, S);
apply ->
i_apply(E, Ctxt, Ren, Env, S);
call ->
i_call(E, Ctxt, Ren, Env, S);
primop ->
i_primop(E, Ren, Env, S);
'try' ->
i_try(E, Ctxt, Ren, Env, S);
'catch' ->
i_catch(E, Ctxt, Ren, Env, S);
binary ->
i_binary(E, Ren, Env, S);
module ->
i_module(E, Ctxt, Ren, Env, S)
end
end.
i_data(E, Ctxt, Ren, Env, S) ->
case is_literal(E) of
true ->
%% This is the `(const c)' case of the original algorithm:
%% literal terms which (regardless of size) do not need to
%% be constructed dynamically at runtime - boldly assuming
%% that the compiler/runtime system can handle this.
case Ctxt of
effect ->
%% Reduce useless constants to a simple value.
{void(), count_size(weight(literal), S)};
_ ->
%% (In Erlang, we cannot set all non-`false'
%% constants to `true' in a `test' context, like we
%% could do in Lisp or C, so the above is the only
%% special case to be handled here.)
{E, count_size(weight(literal), S)}
end;
false ->
%% Data constructors are like to calls to safe built-in
%% functions, for which we can "decide to inline"
%% immediately; there is no need to create operand
%% structures. In `effect' context, we can simply make a
%% sequence of the argument expressions, also visited in
%% `effect' context. In all other cases, the arguments are
%% visited for value.
case Ctxt of
effect ->
%% Note that this will count the sizes of the
%% subexpressions, even though some or all of them
%% might be discarded by the sequencing afterwards.
{Es1, S1} = mapfoldl(fun (E, S) ->
i(E, effect, Ren, Env,
S)
end,
S, data_es(E)),
E1 = foldl(fun (E1, E2) -> make_seq(E1, E2) end,
void(), Es1),
{E1, S1};
_ ->
{Es1, S1} = mapfoldl(fun (E, S) ->
i(E, value, Ren, Env,
S)
end,
S, data_es(E)),
%% The total size/cost is the base cost for a data
%% constructor plus the cost for storing each
%% element.
N = weight(data) + length(Es1) * weight(element),
S2 = count_size(N, S1),
{update_data(E, data_type(E), Es1), S2}
end
end.
%% This is the `(ref x)' (variable use) case of the original algorithm.
%% Note that binding occurrences are always handled in the respective
%% cases of the binding constructs.
i_var(E, Ctxt, Ren, Env, S) ->
case Ctxt of
effect ->
%% Reduce useless variable references to a simple constant.
%% This also avoids useless visiting of bound operands.
{void(), count_size(weight(literal), S)};
_ ->
Name = var_name(E),
case env__lookup(ren__map(Name, Ren), Env) of
{ok, R} ->
case R#ref.opnd of
undefined ->
%% The variable is not associated with an
%% argument expression; just residualize it.
residualize_var(R, S);
Opnd ->
i_var_1(R, Opnd, Ctxt, Env, S)
end;
error ->
%% The variable is unbound. (It has not been
%% accidentally captured, however, or it would have
%% been in the environment.) We leave it as it is,
%% without any warning.
{E, count_size(weight(var), S)}
end
end.
%% This first visits the bound operand and then does copy propagation.
%% Note that we must first set the "inner-pending" flag, and clear the
%% flag afterwards.
i_var_1(R, Opnd, Ctxt, Env, S) ->
%% If the operand is already "inner-pending", it is residualised.
%% (In Lisp/C, if the variable might be assigned to, it should also
%% be residualised.)
L = Opnd#opnd.loc,
case st__test_inner_pending(L, S) of
true ->
residualize_var(R, S);
false ->
S1 = st__mark_inner_pending(L, S),
case catch {ok, visit(Opnd, S1)} of
{ok, {E, S2}} ->
%% Note that we pass the current environment and
%% context to `copy', but not the current renaming.
S3 = st__clear_inner_pending(L, S2),
copy(R, Opnd, E, Ctxt, Env, S3);
{'EXIT', X} ->
exit(X);
X ->
%% If we use destructive update for the
%% `inner-pending' flag, we must make sure to clear
%% it also if we make a nonlocal return.
st__clear_inner_pending(Opnd#opnd.loc, S1),
throw(X)
end
end.
%% A multiple-value aggregate `<e1, ..., en>'. This is very much like a
%% tuple data constructor `{e1, ..., en}'; cf. `i_data' for details.
i_values(E, Ctxt, Ren, Env, S) ->
case values_es(E) of
[E1] ->
%% Single-value aggregates can be dropped; they are simply
%% notation.
i(E1, Ctxt, Ren, Env, S);
Es ->
%% In `effect' context, we can simply make a sequence of the
%% argument expressions, also visited in `effect' context.
%% In all other cases, the arguments are visited for value.
case Ctxt of
effect ->
{Es1, S1} =
mapfoldl(fun (E, S) ->
i(E, effect, Ren, Env, S)
end,
S, Es),
E1 = foldl(fun (E1, E2) ->
make_seq(E1, E2)
end,
void(), Es1),
{E1, S1}; % drop annotations on E
_ ->
{Es1, S1} = mapfoldl(fun (E, S) ->
i(E, value, Ren, Env,
S)
end,
S, Es),
%% Aggregating values does not write them to memory,
%% so we count no extra cost per element.
S2 = count_size(weight(values), S1),
{update_c_values(E, Es1), S2}
end
end.
%% A let-expression `let <v1,...,vn> = e0 in e1' is semantically
%% equivalent to a case-expression `case e0 of <v1,...,vn> when 'true'
%% -> e1 end'. As a special case, `let <v> = e0 in e1' is also
%% equivalent to `apply fun (v) -> e0 (e1)'. However, for efficiency,
%% and in order to allow the handling of `case' clauses to introduce new
%% let-expressions without entering an infinite rewrite loop, we handle
%% these directly.
%%% %% Rewriting a `let' to an equivalent expression.
%%% i_let(E, Ctxt, Ren, Env, S) ->
%%% case let_vars(E) of
%%% [V] ->
%%% E1 = update_c_apply(E, c_fun([V], let_body(E)), [let_arg(E)]),
%%% i(E1, Ctxt, Ren, Env, S);
%%% Vs ->
%%% C = c_clause(Vs, abstract(true), let_body(E)),
%%% E1 = update_c_case(E, let_arg(E), [C]),
%%% i(E1, Ctxt, Ren, Env, S)
%%% end.
i_let(E, Ctxt, Ren, Env, S) ->
case let_vars(E) of
[V] ->
i_let_1(V, E, Ctxt, Ren, Env, S);
Vs ->
%% Visit the argument expression in `value' context, to
%% simplify it as far as possible.
{A, S1} = i(let_arg(E), value, Ren, Env, S),
case get_components(length(Vs), result(A)) of
{true, As} ->
%% Note that only the components of the result of
%% `A' are passed on; any effects are hoisted.
{E1, S2} = i_let_2(Vs, As, E, Ctxt, Ren, Env, S1),
{hoist_effects(A, E1), S2};
false ->
%% We cannot do anything with this `let', since the
%% variables cannot be matched against the argument
%% components. Just visit the variables for renaming
%% and visit the body for value (cf. `i_fun').
{_, Ren1, Env1, S2} = bind_locals(Vs, Ren, Env, S1),
Vs1 = i_params(Vs, Ren1, Env1),
%% The body is always visited for value here.
{B, S3} = i(let_body(E), value, Ren1, Env1, S2),
S4 = count_size(weight('let'), S3),
{update_c_let(E, Vs1, A, B), S4}
end
end.
%% Single-variable `let' binding.
i_let_1(V, E, Ctxt, Ren, Env, S) ->
%% Make an operand structure for the argument expression, create a
%% local binding from the parameter to the operand structure, and
%% visit the body. Finally create necessary bindings and/or set
%% flags.
{Opnd, S1} = make_opnd(let_arg(E), Ren, Env, S),
{[R], Ren1, Env1, S2} = bind_locals([V], [Opnd], Ren, Env, S1),
{E1, S3} = i(let_body(E), Ctxt, Ren1, Env1, S2),
i_let_3([R], [Opnd], E1, S3).
%% Multi-variable `let' binding.
i_let_2(Vs, As, E, Ctxt, Ren, Env, S) ->
%% Make operand structures for the argument components. Note that
%% since the argument has already been visited at this point, we use
%% the identity renaming for the operands.
{Opnds, S1} = mapfoldl(fun (E, S) ->
make_opnd(E, ren__identity(), Env, S)
end,
S, As),
%% Create local bindings from the parameters to their respective
%% operand structures, and visit the body.
{Rs, Ren1, Env1, S2} = bind_locals(Vs, Opnds, Ren, Env, S1),
{E1, S3} = i(let_body(E), Ctxt, Ren1, Env1, S2),
i_let_3(Rs, Opnds, E1, S3).
i_let_3(Rs, Opnds, E, S) ->
%% Create necessary bindings and/or set flags.
{E1, S1} = make_let_bindings(Rs, E, S),
%% We must also create evaluation for effect, for any unused
%% operands, as after an application expression.
residualize_operands(Opnds, E1, S1).
%% A sequence `do e1 e2', written `(seq e1 e2)' in the original
%% algorithm, where `e1' is evaluated for effect only (since its value
%% is not used), and `e2' yields the final value. Note that we use
%% `make_seq' to recompose the sequence after visiting the parts.
i_seq(E, Ctxt, Ren, Env, S) ->
{E1, S1} = i(seq_arg(E), effect, Ren, Env, S),
{E2, S2} = i(seq_body(E), Ctxt, Ren, Env, S1),
%% A sequence has no cost in itself.
{make_seq(E1, E2), S2}.
%% The `case' switch of Core Erlang is rather different from the boolean
%% `(if e1 e2 e3)' case of the original algorithm, but the central idea
%% is the same: if, given the simplified switch expression (which is
%% visited in `value' context - a boolean `test' context would not be
%% generally useful), there is a clause which could definitely be
%% selected, such that no clause before it can possibly be selected,
%% then we can eliminate all other clauses. (And even if this is not the
%% case, some clauses can often be eliminated.) Furthermore, if a clause
%% can be selected, we can replace the case-expression (including the
%% switch expression) with the body of the clause and a set of zero or
%% more let-bindings of subexpressions of the switch expression. (In the
%% simplest case, the switch expression is evaluated only for effect.)
i_case(E, Ctxt, Ren, Env, S) ->
%% First visit the switch expression in `value' context, to simplify
%% it as far as possible. Note that only the result part is passed
%% on to the clause matching below; any effects are hoisted.
{A, S1} = i(case_arg(E), value, Ren, Env, S),
A1 = result(A),
%% Propagating an application context into the branches could cause
%% the arguments of the application to be evaluated *after* the
%% switch expression, but *before* the body of the selected clause.
%% Such interleaving is not allowed in general, and it does not seem
%% worthwile to make a more powerful transformation here. Therefore,
%% the clause bodies are conservatively visited for value if the
%% context is `application'.
Ctxt1 = safe_context(Ctxt),
{E1, S2} = case get_components(case_arity(E), A1) of
{true, As} ->
i_case_1(As, E, Ctxt1, Ren, Env, S1);
false ->
i_case_1([], E, Ctxt1, Ren, Env, S1)
end,
{hoist_effects(A, E1), S2}.
i_case_1(As, E, Ctxt, Ren, Env, S) ->
case i_clauses(As, case_clauses(E), Ctxt, Ren, Env, S) of
{false, {As1, Vs, Env1, Cs}, S1} ->
%% We still have a list of clauses. Sanity check:
if Cs == [] ->
report_warning("empty list of clauses "
"in residual program!.\n");
true ->
ok
end,
{A, S2} = i(c_values(As1), value, ren__identity(), Env1,
S1),
{E1, S3} = i_case_2(Cs, A, E, S2),
i_case_3(Vs, Env1, E1, S3);
{true, {_, Vs, Env1, [C]}, S1} ->
%% A single clause was selected; we just take the body.
i_case_3(Vs, Env1, clause_body(C), S1)
end.
%% Check if all clause bodies are actually equivalent expressions that
%% do not depent on pattern variables (this sometimes occurs as a
%% consequence of inlining, e.g., all branches might yield 'true'), and
%% if so, replace the `case' with a sequence, first evaluating the
%% clause selection for effect, then evaluating one of the clause bodies
%% for its value. (Unless the switch contains a catch-all clause, the
%% clause selection must be evaluated for effect, since there is no
%% guarantee that any of the clauses will actually match. Assuming that
%% some clause always matches could make an undefined program produce a
%% value.) This makes the final size less than what was accounted for
%% when visiting the clauses, but currently we don't try to adjust for
%% this.
i_case_2(Cs, A, E, S) ->
case equivalent_clauses(Cs) of
false ->
%% Count the base sizes for the remaining clauses; pattern
%% and guard sizes are already counted.
N = weight('case') + weight(clause) * length(Cs),
S1 = count_size(N, S),
{update_c_case(E, A, Cs), S1};
true ->
case cerl_clauses:any_catchall(Cs) of
true ->
%% We know that some clause must be selected, so we
%% can drop all the testing as well.
E1 = make_seq(A, clause_body(hd(Cs))),
{E1, S};
false ->
%% The clause selection must be performed for
%% effect.
E1 = update_c_case(E, A,
set_clause_bodies(Cs, void())),
{make_seq(E1, clause_body(hd(Cs))), S}
end
end.
i_case_3(Vs, Env, E, S) ->
%% For the variables bound to the switch expression subexpressions,
%% make let bindings or create evaluation for effect.
Rs = [env__get(var_name(V), Env) || V <- Vs],
{E1, S1} = make_let_bindings(Rs, E, S),
Opnds = [R#ref.opnd || R <- Rs],
residualize_operands(Opnds, E1, S1).
%% This function takes a sequence of switch expressions `Es' (which can
%% be the empty list if these are unknown) and a list `Cs' of clauses,
%% and returns `{Match, {As, Vs, Env1, Cs1}, S1}' where `As' is a list
%% of residual switch expressions, `Vs' the list of variables used in
%% the templates, `Env1' the environment for the templates, and `Cs1'
%% the list of residual clauses. `Match' is `true' if some clause could
%% be shown to definitely match (in this case, `Cs1' contains exactly
%% one element), and `false' otherwise. `S1' is the new state. The given
%% `Ctxt' is the context to be used for visiting the body of clauses.
%%
%% Visiting a clause basically amounts to extending the environment for
%% all variables in the pattern, as for a `fun' (cf. `i_fun'),
%% propagating match information if possible, and visiting the guard and
%% body in the new environment.
%%
%% To make it cheaper to do handle a set of clauses, and to avoid
%% unnecessarily exceeding the size limit, we avoid visiting the bodies
%% of clauses which are subsequently removed, by dividing the visiting
%% of a clause into two stages: first construct the environment(s) and
%% visit the pattern (for renaming) and the guard (for value), then
%% reduce the switch as much as possible, and lastly visit the body.
i_clauses(Cs, Ctxt, Ren, Env, S) ->
i_clauses([], Cs, Ctxt, Ren, Env, S).
i_clauses(Es, Cs, Ctxt, Ren, Env, S) ->
%% Create templates for the switch expressions.
{Ts, {Vs, Env0}} = mapfoldl(fun (E, {Vs, Env}) ->
{T, Vs1, Env1} =
make_template(E, Env),
{T, {Vs1 ++ Vs, Env1}}
end,
{[], Env}, Es),
%% Make operand structures for the switch subexpression templates
%% (found in `Env0') and add proper ref-structure bindings to the
%% environment. Since the subexpressions in general can be
%% interdependent (Vs is in reverse-dependency order), the
%% environment (and renaming) must be created incrementally. Note
%% that since the switch expressions have been visited already, the
%% identity renaming is used for the operands.
Vs1 = lists:reverse(Vs),
{Ren1, Env1, S1} =
foldl(fun (V, {Ren, Env, S}) ->
E = env__get(var_name(V), Env0),
{Opnd, S_1} = make_opnd(E, ren__identity(), Env,
S),
{_, Ren1, Env1, S_2} = bind_locals([V], [Opnd],
Ren, Env, S_1),
{Ren1, Env1, S_2}
end,
{Ren, Env, S}, Vs1),
%% First we visit the head of each individual clause, renaming
%% pattern variables, inserting let-bindings in the guard and body,
%% and visiting the guard. The information used for visiting the
%% clause body will be prefixed to the clause annotations.
{Cs1, S2} = mapfoldl(fun (C, S) ->
i_clause_head(C, Ts, Ren1, Env1, S)
end,
S1, Cs),
%% Now that the clause guards have been reduced as far as possible,
%% we can attempt to reduce the clauses.
As = [hd(get_ann(T)) || T <- Ts],
case cerl_clauses:reduce(Cs1, Ts) of
{false, Cs2} ->
%% We still have one or more clauses (with associated
%% extended environments). Their bodies have not yet been
%% visited, so we do that (in the respective safe
%% environments, adding the sizes of the visited heads to
%% the current size counter) and return the final list of
%% clauses.
{Cs3, S3} = mapfoldl(
fun (C, S) ->
i_clause_body(C, Ctxt, S)
end,
S2, Cs2),
{false, {As, Vs1, Env1, Cs3}, S3};
{true, {C, _}} ->
%% A clause C could be selected (the bindings have already
%% been added to the guard/body). Note that since the clause
%% head will probably be discarded, its size is not counted.
{C1, Ren2, Env2, _} = get_clause_extras(C),
{B, S3} = i(clause_body(C), Ctxt, Ren2, Env2, S2),
C2 = update_c_clause(C1, clause_pats(C1), clause_guard(C1), B),
{true, {As, Vs1, Env1, [C2]}, S3}
end.
%% This visits the head of a clause, renames pattern variables, inserts
%% let-bindings in the guard and body, and does inlining on the guard
%% expression. Returns a list of pairs `{NewClause, Data}', where `Data'
%% is `{Renaming, Environment, Size}' used for visiting the body of the
%% new clause.
i_clause_head(C, Ts, Ren, Env, S) ->
%% Match the templates against the (non-renamed) patterns to get the
%% available information about matching subexpressions. We don't
%% care at this point whether an exact match/nomatch is detected.
Ps = clause_pats(C),
Bs = case cerl_clauses:match_list(Ps, Ts) of
{_, Bs1} -> Bs1;
none -> []
end,
%% The patterns must be visited for renaming; cf. `i_pattern'. We
%% use a passive size counter for visiting the patterns and the
%% guard (cf. `visit'), because we do not know at this stage whether
%% the clause will be kept or not; the final value of the counter is
%% included in the returned value below.
{_, Ren1, Env1, S1} = bind_locals(clause_vars(C), Ren, Env, S),
S2 = new_passive_size(get_size_limit(S1), S1),
{Ps1, S3} = mapfoldl(fun (P, S) ->
i_pattern(P, Ren1, Env1, Ren, Env, S)
end,
S2, Ps),
%% Rewrite guard and body and visit the guard for value. Discard the
%% latter size count if the guard turns out to be a constant.
G = add_match_bindings(Bs, clause_guard(C)),
B = add_match_bindings(Bs, clause_body(C)),
{G1, S4} = i(G, value, Ren1, Env1, S3),
S5 = case is_literal(G1) of
true ->
revert_size(S3, S4);
false ->
S4
end,
%% Revert to the size counter we had on entry to this function. The
%% environment and renaming, together with the size of the clause
%% head, are prefixed to the annotations for later use.
Size = get_size_value(S5),
C1 = update_c_clause(C, Ps1, G1, B),
{set_clause_extras(C1, Ren1, Env1, Size), revert_size(S, S5)}.
add_match_bindings(Bs, E) ->
%% Don't waste time if the variables definitely cannot be used.
%% (Most guards are simply `true'.)
case is_literal(E) of
true ->
E;
false ->
Vs = [V || {V, E} <- Bs, E /= any],
Es = [hd(get_ann(E)) || {_V, E} <- Bs, E /= any],
c_let(Vs, c_values(Es), E)
end.
i_clause_body(C0, Ctxt, S) ->
{C, Ren, Env, Size} = get_clause_extras(C0),
S1 = count_size(Size, S),
{B, S2} = i(clause_body(C), Ctxt, Ren, Env, S1),
C1 = update_c_clause(C, clause_pats(C), clause_guard(C), B),
{C1, S2}.
get_clause_extras(C) ->
[{Ren, Env, Size} | As] = get_ann(C),
{set_ann(C, As), Ren, Env, Size}.
set_clause_extras(C, Ren, Env, Size) ->
As = [{Ren, Env, Size} | get_ann(C)],
set_ann(C, As).
%% This is the `(lambda x e)' case of the original algorithm. A
%% `fun' is like a lambda expression, but with a varying number of
%% parameters; possibly zero.
i_fun(E, Ctxt, Ren, Env, S) ->
case Ctxt of
effect ->
%% Reduce useless `fun' expressions to a simple constant;
%% visiting the body would be a waste of time, and could
%% needlessly mark variables as referenced.
{void(), count_size(weight(literal), S)};
value ->
%% Note that the variables are visited as patterns.
Vs = fun_vars(E),
{_, Ren1, Env1, S1} = bind_locals(Vs, Ren, Env, S),
Vs1 = i_params(Vs, Ren1, Env1),
%% The body is always visited for value.
{B, S2} = i(fun_body(E), value, Ren1, Env1, S1),
%% We don't bother to include the exact number of free
%% variables in the cost for creating a fun-value.
S3 = count_size(weight('fun'), S2),
%% Inlining might have duplicated code, so we must remove
%% any 'id'-annotations from the original fun-expression.
%% (This forces a later stage to invent new id:s.) This is
%% necessary as long as fun:s may still need to be
%% identified the old way. Function variables that are not
%% in application context also have such annotations, but
%% the inlining will currently lose all annotations on
%% variable references (I think), so that's not a problem.
{set_ann(c_fun(Vs1, B), kill_id_anns(get_ann(E))), S3};
#app{} ->
%% An application of a fun-expression (in the source code)
%% is handled by going directly to `inline'; this is never
%% residualised, and we don't set up new counters here. Note
%% that inlining of copy-propagated fun-expressions is done
%% in `copy'; not here.
inline(E, Ctxt, Ren, Env, S)
end.
%% A `letrec' requires a circular environment, but is otherwise like a
%% `let', i.e. like a direct lambda application. Note that only
%% fun-expressions (lambda abstractions) may occur in the right-hand
%% side of each definition.
i_letrec(E, Ctxt, Ren, Env, S) ->
%% Note that we pass an empty list for the auto-referenced
%% (exported) functions here.
{Es, B, _, S1} = i_letrec(letrec_defs(E), letrec_body(E), [], Ctxt,
Ren, Env, S),
%% If no bindings remain, only the body is returned.
case Es of
[] ->
{B, S1}; % drop annotations on E
_ ->
S2 = count_size(weight(letrec), S1),
{update_c_letrec(E, Es, B), S2}
end.
%% The major part of this is shared by letrec-expressions and module
%% definitions alike.
i_letrec(Es, B, Xs, Ctxt, Ren, Env, S) ->
%% First, we create operands with dummy renamings and environments,
%% and with fresh store locations for cached expressions and operand
%% info.
{Opnds, S1} = mapfoldl(fun ({_, E}, S) ->
make_opnd(E, undefined, undefined, S)
end,
S, Es),
%% Then we make recursive bindings for the definitions.
{Rs, Ren1, Env1, S2} = bind_recursive([F || {F, _} <- Es],
Opnds, Ren, Env, S1),
%% For the function variables listed in Xs (none for a
%% letrec-expression), we must make sure that the corresponding
%% operand expressions are visited and that the definitions are
%% marked as referenced; we also need to return the possibly renamed
%% function variables.
{Xs1, S3} =
mapfoldl(
fun (X, S) ->
Name = ren__map(var_name(X), Ren1),
case env__lookup(Name, Env1) of
{ok, R} ->
S_1 = i_letrec_export(R, S),
{ref_to_var(R), S_1};
error ->
%% We just skip any exports that are not
%% actually defined here, and generate a
%% warning message.
{N, A} = var_name(X),
report_warning("export `~w'/~w "
"not defined.\n", [N, A]),
{X, S}
end
end,
S2, Xs),
%% At last, we can then visit the body.
{B1, S4} = i(B, Ctxt, Ren1, Env1, S3),
%% Finally, we create new letrec-bindings for any and all
%% residualised definitions. All referenced functions should have
%% been visited; the call to `visit' below is expected to retreive a
%% cached expression.
Rs1 = keep_referenced(Rs, S4),
{Es1, S5} = mapfoldl(fun (R, S) ->
{E_1, S_1} = visit(R#ref.opnd, S),
{{ref_to_var(R), E_1}, S_1}
end,
S4, Rs1),
{Es1, B1, Xs1, S5}.
%% This visits the operand for a function definition exported by a
%% `letrec' (which is really a `module' module definition, since normal
%% letrecs have no export declarations). Only the updated state is
%% returned. We must handle the "inner-pending" flag when doing this;
%% cf. `i_var'.
i_letrec_export(R, S) ->
Opnd = R#ref.opnd,
S1 = st__mark_inner_pending(Opnd#opnd.loc, S),
{_, S2} = visit(Opnd, S1),
{_, S3} = residualize_var(R, st__clear_inner_pending(Opnd#opnd.loc,
S2)),
S3.
%% This is the `(call e1 e2)' case of the original algorithm. The only
%% difference is that we must handle multiple (or no) operand
%% expressions.
i_apply(E, Ctxt, Ren, Env, S) ->
{Opnds, S1} = mapfoldl(fun (E, S) ->
make_opnd(E, Ren, Env, S)
end,
S, apply_args(E)),
%% Allocate a new app-context location and set up an application
%% context structure containing the surrounding context.
{L, S2} = st__new_app_loc(S1),
Ctxt1 = #app{opnds = Opnds, ctxt = Ctxt, loc = L},
%% Visit the operator expression in the new call context.
{E1, S3} = i(apply_op(E), Ctxt1, Ren, Env, S2),
%% Check the "inlined" flag to find out what to do next. (The store
%% location could be recycled after the flag has been tested, but
%% there is no real advantage to that, because in practice, only
%% 4-5% of all created store locations will ever be reused, while
%% there will be a noticable overhead for managing the free list.)
case st__get_app_inlined(L, S3) of
true ->
%% The application was inlined, so we have the final
%% expression in `E1'. We just have to handle any operands
%% that need to be residualized for effect only (i.e., those
%% the values of which are not used).
residualize_operands(Opnds, E1, S3);
false ->
%% Otherwise, `E1' is the residual operator expression. We
%% make sure all operands are visited, and rebuild the
%% application.
{Es, S4} = mapfoldl(fun (Opnd, S) ->
visit_and_count_size(Opnd, S)
end,
S3, Opnds),
N = apply_size(length(Es)),
{update_c_apply(E, E1, Es), count_size(N, S4)}
end.
apply_size(A) ->
weight(apply) + weight(argument) * A.
%% Since it is not the task of this transformation to handle
%% cross-module inlining, all inter-module calls are handled by visiting
%% the components (the module and function name, and the arguments of
%% the call) for value. In `effect' context, if the function itself is
%% known to be completely effect free, the call can be discarded and the
%% arguments evaluated for effect. Otherwise, if all the visited
%% arguments are to constants, and the function is known to be safe to
%% execute at compile time, then we try to evaluate the call. If
%% evaluation completes normally, the call is replaced by the result;
%% otherwise the call is residualised.
i_call(E, Ctxt, Ren, Env, S) ->
{M, S1} = i(call_module(E), value, Ren, Env, S),
{F, S2} = i(call_name(E), value, Ren, Env, S1),
As = call_args(E),
Arity = length(As),
%% Check if the name of the called function is static. If so,
%% discard the size counts performed above, since the values will
%% not cause any runtime cost.
Static = is_c_atom(M) and is_c_atom(F),
S3 = case Static of
true ->
revert_size(S, S2);
false ->
S2
end,
case Ctxt of
effect when Static == true ->
case is_safe_call(atom_val(M), atom_val(F), Arity) of
true ->
%% The result will not be used, and the call is
%% effect free, so we create a multiple-value
%% aggregate containing the (not yet visited)
%% arguments and process that instead.
i(c_values(As), effect, Ren, Env, S3);
false ->
%% We are not allowed to simply discard the call,
%% but we can try to evaluate it.
i_call_1(Static, M, F, Arity, As, E, Ctxt, Ren, Env,
S3)
end;
_ ->
i_call_1(Static, M, F, Arity, As, E, Ctxt, Ren, Env, S3)
end.
i_call_1(Static, M, F, Arity, As, E, Ctxt, Ren, Env, S) ->
%% Visit the arguments for value.
{As1, S1} = mapfoldl(fun (X, A) -> i(X, value, Ren, Env, A) end,
S, As),
case Static of
true ->
case erl_bifs:is_pure(atom_val(M), atom_val(F), Arity) of
true ->
%% It is allowed to evaluate this at compile time.
case all_static(As1) of
true ->
i_call_3(M, F, As1, E, Ctxt, Env, S1);
false ->
%% See if the call can be rewritten instead.
i_call_4(M, F, As1, E, Ctxt, Env, S1)
end;
false ->
i_call_2(M, F, As1, E, S1)
end;
false ->
i_call_2(M, F, As1, E, S1)
end.
%% Residualise the call.
i_call_2(M, F, As, E, S) ->
N = weight(call) + weight(argument) * length(As),
{update_c_call(E, M, F, As), count_size(N, S)}.
%% Attempt to evaluate the call to yield a literal; if that fails, try
%% to rewrite the expression.
i_call_3(M, F, As, E, Ctxt, Env, S) ->
%% Note that we extract the results of argument expessions here; the
%% expressions could still be sequences with side effects.
Vs = [concrete(result(A)) || A <- As],
case catch {ok, apply(atom_val(M), atom_val(F), Vs)} of
{ok, V} ->
%% Evaluation completed normally - try to turn the result
%% back into a syntax tree (representing a literal).
case is_literal_term(V) of
true ->
%% Make a sequence of the arguments (as a
%% multiple-value aggregate) and the final value.
S1 = count_size(weight(values), S),
S2 = count_size(weight(literal), S1),
{make_seq(c_values(As), abstract(V)), S2};
false ->
%% The result could not be represented as a literal.
i_call_4(M, F, As, E, Ctxt, Env, S)
end;
_ ->
%% The evaluation attempt did not complete normally.
i_call_4(M, F, As, E, Ctxt, Env, S)
end.
%% Rewrite the expression, if possible, otherwise residualise it.
i_call_4(M, F, As, E, Ctxt, Env, S) ->
case reduce_bif_call(atom_val(M), atom_val(F), As, Env) of
false ->
%% Nothing more to be done - residualise the call.
i_call_2(M, F, As, E, S);
{true, E1} ->
%% We revisit the result, because the rewriting might have
%% opened possibilities for further inlining. Since the
%% parts have already been visited once, we use the identity
%% renaming here.
i(E1, Ctxt, ren__identity(), Env, S)
end.
%% For now, we assume that primops cannot be evaluated at compile time,
%% probably being too special. Also, we have no knowledge about their
%% side effects.
i_primop(E, Ren, Env, S) ->
%% Visit the arguments for value.
{As, S1} = mapfoldl(fun (E, S) ->
i(E, value, Ren, Env, S)
end,
S, primop_args(E)),
N = weight(primop) + weight(argument) * length(As),
{update_c_primop(E, primop_name(E), As), count_size(N, S1)}.
%% This is like having an expression with an extra fun-expression
%% attached for "exceptional cases"; actually, there are exactly two
%% parameter variables for the body, but they are easiest handled as if
%% their number might vary, just as for a `fun'.
i_try(E, Ctxt, Ren, Env, S) ->
%% The argument expression is evaluated in `value' context, and the
%% surrounding context is propagated into both branches. We do not
%% try to recognize cases when the protected expression will
%% actually raise an exception. Note that the variables are visited
%% as patterns.
{A, S1} = i(try_arg(E), value, Ren, Env, S),
Vs = try_vars(E),
{_, Ren1, Env1, S2} = bind_locals(Vs, Ren, Env, S1),
Vs1 = i_params(Vs, Ren1, Env1),
{B, S3} = i(try_body(E), Ctxt, Ren1, Env1, S2),
case is_safe(A) of
true ->
%% The `try' wrapper can be dropped in this case. Since the
%% expressions have been visited already, the identity
%% renaming is used when we revisit the new let-expression.
i(c_let(Vs1, A, B), Ctxt, ren__identity(), Env, S3);
false ->
Evs = try_evars(E),
{_, Ren2, Env2, S4} = bind_locals(Evs, Ren, Env, S3),
Evs1 = i_params(Evs, Ren2, Env2),
{H, S5} = i(try_handler(E), Ctxt, Ren2, Env2, S4),
S6 = count_size(weight('try'), S5),
{update_c_try(E, A, Vs1, B, Evs1, H), S6}
end.
%% A special case of try-expressions:
i_catch(E, Ctxt, Ren, Env, S) ->
%% We cannot propagate application contexts into the catch.
{E1, S1} = i(catch_body(E), safe_context(Ctxt), Ren, Env, S),
case is_safe(E1) of
true ->
%% The `catch' wrapper can be dropped in this case.
{E1, S1};
false ->
S2 = count_size(weight('catch'), S1),
{update_c_catch(E, E1), S2}
end.
%% A receive-expression is very much like a case-expression, with the
%% difference that we do not have access to a switch expression, since
%% the value being switched on is taken from the mailbox. The fact that
%% the receive-expression may iterate over an arbitrary number of
%% messages is not of interest to us. All we can do here is to visit its
%% subexpressions, and possibly eliminate definitely unselectable
%% clauses.
i_receive(E, Ctxt, Ren, Env, S) ->
%% We first visit the expiry expression (for value) and the expiry
%% body (in the surrounding context).
{T, S1} = i(receive_timeout(E), value, Ren, Env, S),
{B, S2} = i(receive_action(E), Ctxt, Ren, Env, S1),
%% Then we visit the clauses. Note that application contexts may not
%% in general be propagated into the branches (and the expiry body),
%% because the execution of the `receive' may remove a message from
%% the mailbox as a side effect; the situation is thus analogous to
%% that in a `case' expression.
Ctxt1 = safe_context(Ctxt),
case i_clauses(receive_clauses(E), Ctxt1, Ren, Env, S2) of
{false, {[], _, _, Cs}, S3} ->
%% We still have a list of clauses. If the list is empty,
%% and the expiry expression is the integer zero, the
%% expression reduces to the expiry body.
if Cs == [] ->
case is_c_int(T) andalso (int_val(T) == 0) of
true ->
{B, S3};
false ->
i_receive_1(E, Cs, T, B, S3)
end;
true ->
i_receive_1(E, Cs, T, B, S3)
end;
{true, {_, _, _, Cs}, S3} ->
%% Cs is a single clause that will always be matched (if a
%% message exists), but we must keep the `receive' statement
%% in order to fetch the message from the mailbox.
i_receive_1(E, Cs, T, B, S3)
end.
i_receive_1(E, Cs, T, B, S) ->
%% Here, we just add the base sizes for the receive-expression
%% itself and for each remaining clause; cf. `case'.
N = weight('receive') + weight(clause) * length(Cs),
{update_c_receive(E, Cs, T, B), count_size(N, S)}.
%% A module definition is like a `letrec', with some add-ons (export and
%% attribute declarations) but without an explicit body. Actually, the
%% exporting of function names has the same effect as if there was a
%% body consisting of the list of references to the exported functions.
%% Thus, the exported functions are exactly those which can be
%% referenced from outside the module.
i_module(E, Ctxt, Ren, Env, S) ->
%% Cf. `i_letrec'. Note that we pass a dummy constant value for the
%% "body" parameter.
{Es, _, Xs1, S1} = i_letrec(module_defs(E), void(),
module_exports(E), Ctxt, Ren, Env, S),
%% Sanity check:
case Es of
[] ->
report_warning("no function definitions remaining "
"in module `~s'.\n",
[atom_name(module_name(E))]);
_ ->
ok
end,
E1 = update_c_module(E, module_name(E), Xs1, module_attrs(E), Es),
{E1, count_size(weight(module), S1)}.
%% Binary-syntax expressions are too complicated to do anything
%% interesting with here - that is beyond the scope of this program;
%% also, their construction could have side effects, so even in effect
%% context we can't remove them. (We don't bother to identify cases of
%% "safe" unused binaries which could be removed.)
i_binary(E, Ren, Env, S) ->
%% Visit the segments for value.
{Es, S1} = mapfoldl(fun (E, S) ->
i_bitstr(E, Ren, Env, S)
end,
S, binary_segments(E)),
S2 = count_size(weight(binary), S1),
{update_c_binary(E, Es), S2}.
i_bitstr(E, Ren, Env, S) ->
%% It is not necessary to visit the Unit, Type and Flags fields,
%% since these are always literals.
{Val, S1} = i(bitstr_val(E), value, Ren, Env, S),
{Size, S2} = i(bitstr_size(E), value, Ren, Env, S1),
Unit = bitstr_unit(E),
Type = bitstr_type(E),
Flags = bitstr_flags(E),
S3 = count_size(weight(bitstr), S2),
{update_c_bitstr(E, Val, Size, Unit, Type, Flags), S3}.
%% This is a simplified version of `i_pattern', for lists of parameter
%% variables only. It does not modify the state.
i_params([V | Vs], Ren, Env) ->
Name = ren__map(var_name(V), Ren),
case env__lookup(Name, Env) of
{ok, R} ->
[ref_to_var(R) | i_params(Vs, Ren, Env)];
error ->
report_internal_error("variable `~w' not bound "
"in pattern.\n", [Name]),
exit(error)
end;
i_params([], _, _) ->
[].
%% For ordinary patterns, we just visit to rename variables and count
%% the size/cost. All occurring binding instances of variables should
%% already have been added to the renaming and environment; however, to
%% handle the size expressions of binary-syntax patterns, we must pass
%% the renaming and environment of the containing expression
i_pattern(E, Ren, Env, Ren0, Env0, S) ->
case type(E) of
var ->
%% Count no size.
Name = ren__map(var_name(E), Ren),
case env__lookup(Name, Env) of
{ok, R} ->
{ref_to_var(R), S};
error ->
report_internal_error("variable `~w' not bound "
"in pattern.\n", [Name]),
exit(error)
end;
alias ->
%% Count no size.
V = alias_var(E),
Name = ren__map(var_name(V), Ren),
case env__lookup(Name, Env) of
{ok, R} ->
%% Visit the subpattern and recompose.
V1 = ref_to_var(R),
{P, S1} = i_pattern(alias_pat(E), Ren, Env, Ren0,
Env0, S),
{update_c_alias(E, V1, P), S1};
error ->
report_internal_error("variable `~w' not bound "
"in pattern.\n", [Name]),
exit(error)
end;
binary ->
{Es, S1} = mapfoldl(fun (E, S) ->
i_bitstr_pattern(E, Ren, Env,
Ren0, Env0, S)
end,
S, binary_segments(E)),
S2 = count_size(weight(binary), S1),
{update_c_binary(E, Es), S2};
_ ->
case is_literal(E) of
true ->
{E, count_size(weight(literal), S)};
false ->
{Es1, S1} = mapfoldl(fun (E, S) ->
i_pattern(E, Ren, Env,
Ren0, Env0,
S)
end,
S, data_es(E)),
%% We assume that in general, the elements of the
%% constructor will all be fetched.
N = weight(data) + length(Es1) * weight(element),
S2 = count_size(N, S1),
{update_data(E, data_type(E), Es1), S2}
end
end.
i_bitstr_pattern(E, Ren, Env, Ren0, Env0, S) ->
%% It is not necessary to visit the Unit, Type and Flags fields,
%% since these are always literals. The Value field is a limited
%% pattern - either a literal or an unbound variable. The Size field
%% is a limited expression - either a literal or a variable bound in
%% the environment of the containing expression.
{Val, S1} = i_pattern(bitstr_val(E), Ren, Env, Ren0, Env0, S),
{Size, S2} = i(bitstr_size(E), value, Ren0, Env0, S1),
Unit = bitstr_unit(E),
Type = bitstr_type(E),
Flags = bitstr_flags(E),
S3 = count_size(weight(bitstr), S2),
{update_c_bitstr(E, Val, Size, Unit, Type, Flags), S3}.
%% ---------------------------------------------------------------------
%% Other central inlining functions
%% It is assumed here that `E' is a fun-expression and the context is an
%% app-structure. If the inlining might be aborted for some reason, a
%% corresponding catch should have been set up before entering `inline'.
%%
%% Note: if the inlined body is a lambda abstraction, and the
%% surrounding context of the app-context is also an app-context, the
%% `inlined' flag of the outermost context will be set before that of
%% the inner context is set. E.g.: `let F = fun (X) -> fun (Y) -> E in
%% apply apply F(A)(B)' will propagate the body of F, which is a lambda
%% abstraction, into the outer application context, which will be
%% inlined to produce expression `E', and the flag of the outer context
%% will be set. Upon return, the flag of the inner context will also be
%% set. However, the flags are then tested in innermost-first order.
%% Thus, if some inlining attempt is aborted, the `inlined' flags of any
%% nested app-contexts must be cleared.
%%
%% This implementation does nothing to handle inlining of calls to
%% recursive functions in a smart way. This means that as long as the
%% size and effort counters do not prevent it, the function body will be
%% inlined (i.e., the first iteration will be unrolled), and the
%% recursive calls will be residualized.
inline(E, #app{opnds = Opnds, ctxt = Ctxt, loc = L}, Ren, Env, S) ->
%% Check that the arities match:
Vs = fun_vars(E),
if length(Opnds) /= length(Vs) ->
report_error("function called with wrong number "
"of arguments!\n"),
%% TODO: should really just residualise the call...
exit(error);
true ->
ok
end,
%% Create local bindings for the parameters to their respective
%% operand structures from the app-structure, and visit the body in
%% the context saved in the structure.
{Rs, Ren1, Env1, S1} = bind_locals(Vs, Opnds, Ren, Env, S),
{E1, S2} = i(fun_body(E), Ctxt, Ren1, Env1, S1),
%% Create necessary bindings and/or set flags.
{E2, S3} = make_let_bindings(Rs, E1, S2),
%% Lastly, flag the application as inlined, since the inlining
%% attempt was not aborted before we reached this point.
{E2, st__set_app_inlined(L, S3)}.
%% For the (possibly renamed) argument variables to an inlined call,
%% either create `let' bindings for them, if they are still referenced
%% in the residual expression (in C/Lisp, also if they are assigned to),
%% or otherwise (if they are not referenced or assigned) mark them for
%% evaluation for side effects.
make_let_bindings([R | Rs], E, S) ->
{E1, S1} = make_let_bindings(Rs, E, S),
make_let_binding(R, E1, S1);
make_let_bindings([], E, S) ->
{E, S}.
make_let_binding(R, E, S) ->
%% The `referenced' flag is conservatively computed. We therefore
%% first check some simple cases where parameter R is definitely not
%% referenced in the resulting body E.
case is_literal(E) of
true ->
%% A constant contains no variable references.
make_let_binding_1(R, E, S);
false ->
case is_c_var(E) of
true ->
case var_name(E) =:= R#ref.name of
true ->
%% The body is simply the parameter variable
%% itself. Visit the operand for value and
%% substitute the result for the body.
visit_and_count_size(R#ref.opnd, S);
false ->
%% Not the same variable, so the parameter
%% is not referenced at all.
make_let_binding_1(R, E, S)
end;
false ->
%% Proceed to check the `referenced' flag.
case st__get_var_referenced(R#ref.loc, S) of
true ->
%% The parameter is probably referenced in
%% the residual code (although it might not
%% be). Visit the operand for value and
%% create a let-binding.
{E1, S1} = visit_and_count_size(R#ref.opnd,
S),
S2 = count_size(weight('let'), S1),
{c_let([ref_to_var(R)], E1, E), S2};
false ->
%% The parameter is definitely not
%% referenced.
make_let_binding_1(R, E, S)
end
end
end.
%% This marks the operand for evaluation for effect.
make_let_binding_1(R, E, S) ->
Opnd = R#ref.opnd,
{E, st__set_opnd_effect(Opnd#opnd.loc, S)}.
%% Here, `R' is the ref-structure which is the target of the copy
%% propagation, and `Opnd' is a visited operand structure, to be
%% propagated through `R' if possible - if not, `R' is residualised.
%% `Opnd' is normally the operand that `R' is bound to, and `E' is the
%% result of visiting `Opnd' for value; we pass this as an argument so
%% we don't have to fetch it multiple times (because we don't have
%% constant time access).
%%
%% We also pass the environment of the site of the variable reference,
%% for use when inlining a propagated fun-expression. In the original
%% algorithm by Waddell, the environment used for inlining such cases is
%% the identity mapping, because the fun-expression body has already
%% been visited for value, and their algorithm combines renaming of
%% source-code variables with the looking up of information about
%% residual-code variables. We, however, need to check the environment
%% of the call site when creating new non-shadowed variables, but we
%% must avoid repeated renaming. We therefore separate the renaming and
%% the environment (as in the renaming algorithm of Peyton-Jones and
%% Marlow). This also makes our implementation more general, compared to
%% the original algorithm, because we do not give up on propagating
%% variables that were free in the fun-body.
%%
%% Example:
%%
%% let F = fun (X) -> {'foo', X} in
%% let G = fun (H) -> apply H(F) % F is free in the fun G
%% in apply G(fun (F) -> apply F(42))
%% =>
%% let F = fun (X) -> {'foo', X} in
%% apply (fun (H) -> apply H(F))(fun (F) -> apply F(42))
%% =>
%% let F = fun (X) -> {'foo', X} in
%% apply (fun (F) -> apply F(42))(F)
%% =>
%% let F = fun (X) -> {'foo', X} in
%% apply F(42)
%% =>
%% apply (fun (X) -> {'foo', X})(2)
%% =>
%% {'foo', 42}
%%
%% The original algorithm would give up at stage 4, because F was free
%% in the propagated fun-expression. Our version inlines this example
%% completely.
copy(R, Opnd, E, Ctxt, Env, S) ->
case is_c_var(E) of
true ->
%% The operand reduces to another variable - get its
%% ref-structure and attempt to propagate further.
copy_var(env__get(var_name(E), Opnd#opnd.env), Ctxt, Env,
S);
false ->
%% Apart from variables and functional values (the latter
%% are handled by `copy_1' below), only constant literals
%% are copyable in general; other things, including e.g.
%% tuples `{foo, X}', could cause duplication of work, and
%% are not copy propagated.
case is_literal(E) of
true ->
{E, count_size(weight(literal), S)};
false ->
copy_1(R, Opnd, E, Ctxt, Env, S)
end
end.
copy_var(R, Ctxt, Env, S) ->
%% (In Lisp or C, if this other variable might be assigned to, we
%% should residualize the "parent" instead, so we don't bypass any
%% destructive updates.)
case R#ref.opnd of
undefined ->
%% This variable is not bound to an expression, so just
%% residualize it.
residualize_var(R, S);
Opnd ->
%% Note that because operands are always visited before
%% copied, all copyable operand expressions will be
%% propagated through any number of bindings. If `R' was
%% bound to a constant literal, we would never have reached
%% this point.
case st__lookup_opnd_cache(Opnd#opnd.loc, S) of
error ->
%% The result for this operand is not yet ready
%% (which should mean that it is a recursive
%% reference). Thus, we must residualise the
%% variable.
residualize_var(R, S);
{ok, #cache{expr = E1}} ->
%% The result for the operand is ready, so we can
%% proceed to propagate it.
copy_1(R, Opnd, E1, Ctxt, Env, S)
end
end.
copy_1(R, Opnd, E, Ctxt, Env, S) ->
%% Fun-expression (lambdas) are a bit special; they are copyable,
%% but should preferably not be duplicated, so they should not be
%% copy propagated except into application contexts, where they can
%% be inlined.
case is_c_fun(E) of
true ->
case Ctxt of
#app{} ->
%% First test if the operand is "outer-pending"; if
%% so, don't inline.
case st__test_outer_pending(Opnd#opnd.loc, S) of
false ->
copy_inline(R, Opnd, E, Ctxt, Env, S);
true ->
%% Cyclic reference forced inlining to stop
%% (avoiding infinite unfolding).
residualize_var(R, S)
end;
_ ->
residualize_var(R, S)
end;
false ->
%% We have no other cases to handle here
residualize_var(R, S)
end.
%% This inlines a function value that was propagated to an application
%% context. The inlining is done with an identity renaming (since the
%% expression is already visited) but in the environment of the call
%% site (which is OK because of the no-shadowing strategy for renaming,
%% and because the domain of our environments are the residual-program
%% variables instead of the source-program variables). Note that we must
%% first set the "outer-pending" flag, and clear it afterwards.
copy_inline(R, Opnd, E, Ctxt, Env, S) ->
S1 = st__mark_outer_pending(Opnd#opnd.loc, S),
case catch {ok, copy_inline_1(R, E, Ctxt, Env, S1)} of
{ok, {E1, S2}} ->
{E1, st__clear_outer_pending(Opnd#opnd.loc, S2)};
{'EXIT', X} ->
exit(X);
X ->
%% If we use destructive update for the `outer-pending'
%% flag, we must make sure to clear it upon a nonlocal
%% return.
st__clear_outer_pending(Opnd#opnd.loc, S1),
throw(X)
end.
%% If the current effort counter was passive, we use a new active effort
%% counter with the inherited limit for this particular inlining.
copy_inline_1(R, E, Ctxt, Env, S) ->
case effort_is_active(S) of
true ->
copy_inline_2(R, E, Ctxt, Env, S);
false ->
S1 = new_active_effort(get_effort_limit(S), S),
case catch {ok, copy_inline_2(R, E, Ctxt, Env, S1)} of
{ok, {E1, S2}} ->
%% Revert to the old effort counter.
{E1, revert_effort(S, S2)};
{counter_exceeded, effort, _} ->
%% Aborted this inlining attempt because too much
%% effort was spent. Residualize the variable and
%% revert to the previous state.
residualize_var(R, S);
{'EXIT', X} ->
exit(X);
X ->
throw(X)
end
end.
%% Regardless of whether the current size counter is active or not, we
%% use a new active size counter for each inlining. If the current
%% counter was passive, the new counter gets the inherited size limit;
%% if it was active, the size limit of the new counter will be equal to
%% the remaining budget of the current counter (which itself is not
%% affected by the inlining). This distributes the size budget more
%% evenly over "inlinings within inlinings", so that the whole size
%% budget is not spent on the first few call sites (in an inlined
%% function body) forcing the remaining call sites to be residualised.
copy_inline_2(R, E, Ctxt, Env, S) ->
Limit = case size_is_active(S) of
true ->
get_size_limit(S) - get_size_value(S);
false ->
get_size_limit(S)
end,
%% Add the cost of the application to the new size limit, so we
%% always inline functions that are small enough, even if `Limit' is
%% close to zero at this point. (This is an extension to the
%% original algorithm.)
S1 = new_active_size(Limit + apply_size(length(Ctxt#app.opnds)), S),
case catch {ok, inline(E, Ctxt, ren__identity(), Env, S1)} of
{ok, {E1, S2}} ->
%% Revert to the old size counter.
{E1, revert_size(S, S2)};
{counter_exceeded, size, S2} ->
%% Aborted this inlining attempt because it got too big.
%% Residualize the variable and revert to the old size
%% counter. (It is important that we do not also revert the
%% effort counter here. Because the effort and size counters
%% are always set up together, we know that the effort
%% counter returned in S2 is the same that was passed to
%% `inline'.)
S3 = revert_size(S, S2),
%% If we use destructive update for the `inlined' flag, we
%% must make sure to clear the flags of any nested
%% app-contexts upon aborting; see `inline' for details.
reset_nested_apps(Ctxt, S3), % for effect
residualize_var(R, S3);
{'EXIT', X} ->
exit(X);
X ->
throw(X)
end.
reset_nested_apps(#app{ctxt = Ctxt, loc = L}, S) ->
reset_nested_apps(Ctxt, st__clear_app_inlined(L, S));
reset_nested_apps(_, S) ->
S.
%% ---------------------------------------------------------------------
%% Support functions
new_var(Env) ->
Name = env__new_vname(Env),
c_var(Name).
residualize_var(R, S) ->
S1 = count_size(weight(var), S),
{ref_to_var(R), st__set_var_referenced(R#ref.loc, S1)}.
%% This function returns the value-producing subexpression of any
%% expression. (Except for sequencing expressions, this is the
%% expression itself.)
result(E) ->
case is_c_seq(E) of
true ->
%% Also see `make_seq', which is used in all places to build
%% sequences so that they are always nested in the first
%% position.
seq_body(E);
false ->
E
end.
%% This function rewrites E to `do A1 E' if A is `do A1 A2', and
%% otherwise returns E unchanged.
hoist_effects(A, E) ->
case type(A) of
seq -> make_seq(seq_arg(A), E);
_ -> E
end.
%% This "build sequencing expression" operation assures that sequences
%% are always nested in the first position, which makes it easy to find
%% the actual value-producing expression of a sequence (cf. `result').
make_seq(E1, E2) ->
case is_safe(E1) of
true ->
%% The first expression can safely be dropped.
E2;
false ->
%% If `E1' is a sequence whose final expression has no side
%% effects, then we can lose *that* expression when we
%% compose the new sequence, since its value will not be
%% used.
E3 = case is_c_seq(E1) of
true ->
case is_safe(seq_body(E1)) of
true ->
%% Drop the final expression.
seq_arg(E1);
false ->
E1
end;
false ->
E1
end,
case is_c_seq(E2) of
true ->
%% `E2' is a sequence (E2' E2''), so we must
%% rearrange the nesting to ((E1, E2') E2''), to
%% preserve the invariant. Annotations on `E2' are
%% lost.
c_seq(c_seq(E3, seq_arg(E2)), seq_body(E2));
false ->
c_seq(E3, E2)
end
end.
%% Currently, safe expressions include variables, lambda expressions,
%% constructors with safe subexpressions (this includes atoms, integers,
%% empty lists, etc.), seq-, let- and letrec-expressions with safe
%% subexpressions, try- and catch-expressions with safe subexpressions
%% and calls to safe functions with safe argument subexpressions.
%% Binaries seem too tricky to be considered.
is_safe(E) ->
case is_data(E) of
true ->
is_safe_list(data_es(E));
false ->
case type(E) of
var ->
true;
'fun' ->
true;
values ->
is_safe_list(values_es(E));
'seq' ->
case is_safe(seq_arg(E)) of
true ->
is_safe(seq_body(E));
false ->
false
end;
'let' ->
case is_safe(let_arg(E)) of
true ->
is_safe(let_body(E));
false ->
false
end;
letrec ->
is_safe(letrec_body(E));
'try' ->
%% If the argument expression is not safe, it could
%% be modifying the state; thus, even if the body is
%% safe, the try-expression as a whole would not be.
%% If the argument is safe, the handler is not used.
case is_safe(try_arg(E)) of
true ->
is_safe(try_body(E));
false ->
false
end;
'catch' ->
is_safe(catch_body(E));
call ->
M = call_module(E),
F = call_name(E),
case is_c_atom(M) and is_c_atom(F) of
true ->
As = call_args(E),
case is_safe_list(As) of
true ->
is_safe_call(atom_val(M),
atom_val(F),
length(As));
false ->
false
end;
false ->
false
end;
_ ->
false
end
end.
is_safe_list([E | Es]) ->
case is_safe(E) of
true ->
is_safe_list(Es);
false ->
false
end;
is_safe_list([]) ->
true.
is_safe_call(M, F, A) ->
erl_bifs:is_safe(M, F, A).
%% When setting up local variables, we only create new names if we have
%% to, according to the "no-shadowing" strategy.
make_locals(Vs, Ren, Env) ->
make_locals(Vs, [], Ren, Env).
make_locals([V | Vs], As, Ren, Env) ->
Name = var_name(V),
case env__is_defined(Name, Env) of
false ->
%% The variable need not be renamed. Just make sure that the
%% renaming will map it to itself.
Name1 = Name,
Ren1 = ren__add_identity(Name, Ren);
true ->
%% The variable must be renamed to maintain the no-shadowing
%% invariant. Do the right thing for function variables.
Name1 = case Name of
{A, N} ->
env__new_fname(A, N, Env);
_ ->
env__new_vname(Env)
end,
Ren1 = ren__add(Name, Name1, Ren)
end,
%% This temporary binding is added for correct new-key generation.
Env1 = env__bind(Name1, dummy, Env),
make_locals(Vs, [Name1 | As], Ren1, Env1);
make_locals([], As, Ren, Env) ->
{reverse(As), Ren, Env}.
%% This adds let-bindings for the source code variables in `Es' to the
%% environment `Env'.
%%
%% Note that we always assign a new state location for the
%% residual-program variable, since we cannot know when a location for a
%% particular variable in the source code can be reused.
bind_locals(Vs, Ren, Env, S) ->
Opnds = lists:duplicate(length(Vs), undefined),
bind_locals(Vs, Opnds, Ren, Env, S).
bind_locals(Vs, Opnds, Ren, Env, S) ->
{Ns, Ren1, Env1} = make_locals(Vs, Ren, Env),
{Rs, Env2, S1} = bind_locals_1(Ns, Opnds, [], Env1, S),
{Rs, Ren1, Env2, S1}.
%% Note that the `Vs' are currently not used for anything except the
%% number of variables. If we were maintaining "source-referenced"
%% flags, then the flag in the new variable should be initialized to the
%% current value of the (residual-) referenced-flag of the "parent".
bind_locals_1([N | Ns], [Opnd | Opnds], Rs, Env, S) ->
{R, S1} = new_ref(N, Opnd, S),
Env1 = env__bind(N, R, Env),
bind_locals_1(Ns, Opnds, [R | Rs], Env1, S1);
bind_locals_1([], [], Rs, Env, S) ->
{lists:reverse(Rs), Env, S}.
new_refs(Ns, Opnds, S) ->
new_refs(Ns, Opnds, [], S).
new_refs([N | Ns], [Opnd | Opnds], Rs, S) ->
{R, S1} = new_ref(N, Opnd, S),
new_refs(Ns, Opnds, [R | Rs], S1);
new_refs([], [], Rs, S) ->
{lists:reverse(Rs), S}.
new_ref(N, Opnd, S) ->
{L, S1} = st__new_ref_loc(S),
{#ref{name = N, opnd = Opnd, loc = L}, S1}.
%% This adds recursive bindings for the source code variables in `Es' to
%% the environment `Env'. Note that recursive binding of a set of
%% variables is an atomic operation on the environment - they cannot be
%% added one at a time.
bind_recursive(Vs, Opnds, Ren, Env, S) ->
{Ns, Ren1, Env1} = make_locals(Vs, Ren, Env),
{Rs, S1} = new_refs(Ns, Opnds, S),
%% When this fun-expression is evaluated, it updates the operand
%% structure in the ref-structure to contain the recursively defined
%% environment and the correct renaming.
Fun = fun (R, Env) ->
Opnd = R#ref.opnd,
R#ref{opnd = Opnd#opnd{ren = Ren1, env = Env}}
end,
{Rs, Ren1, env__bind_recursive(Ns, Rs, Fun, Env1), S1}.
safe_context(Ctxt) ->
case Ctxt of
#app{} ->
value;
_ ->
Ctxt
end.
%% Note that the name of a variable encodes its type: a "plain" variable
%% or a function variable. The latter kind also contains an arity number
%% which should be preserved upon renaming.
ref_to_var(#ref{name = Name}) ->
%% If we were maintaining "source-referenced" flags, the annotation
%% `add_ann([#source_ref{loc = L}], E)' should also be done here, to
%% make the algorithm reapplicable. This is however not necessary
%% since there are no destructive variable assignments in Erlang.
c_var(Name).
%% Including the effort counter of the call site assures that the cost
%% of processing an operand via `visit' is charged to the correct
%% counter. In particular, if the effort counter of the call site was
%% passive, the operands will also be processed with a passive counter.
make_opnd(E, Ren, Env, S) ->
{L, S1} = st__new_opnd_loc(S),
C = st__get_effort(S1),
Opnd = #opnd{expr = E, ren = Ren, env = Env, loc = L, effort = C},
{Opnd, S1}.
keep_referenced(Rs, S) ->
[R || R <- Rs, st__get_var_referenced(R#ref.loc, S)].
residualize_operands(Opnds, E, S) ->
foldr(fun (Opnd, {E, S}) -> residualize_operand(Opnd, E, S) end,
{E, S}, Opnds).
%% This is the only case where an operand expression can be visited in
%% `effect' context instead of `value' context.
residualize_operand(Opnd, E, S) ->
case st__get_opnd_effect(Opnd#opnd.loc, S) of
true ->
%% The operand has not been visited, so we do that now, but
%% in `effect' context. (Waddell's algoritm does some stuff
%% here to account specially for the operand size, which
%% appears unnecessary.)
{E1, S1} = i(Opnd#opnd.expr, effect, Opnd#opnd.ren,
Opnd#opnd.env, S),
{make_seq(E1, E), S1};
false ->
{E, S}
end.
%% The `visit' function always visits the operand expression in `value'
%% context (`residualize_operand' visits an unreferenced operand
%% expression in `effect' context when necessary). A new passive size
%% counter is used for visiting the operand, the final value of which is
%% then cached along with the resulting expression.
%%
%% Note that the effort counter of the call site, included in the
%% operand structure, is not a shared object. Thus, the effort budget is
%% actually reused over all occurrences of the operands of a single
%% application. This does not appear to be a problem; just a
%% modification of the algorithm.
visit(Opnd, S) ->
{C, S1} = visit_1(Opnd, S),
{C#cache.expr, S1}.
visit_and_count_size(Opnd, S) ->
{C, S1} = visit_1(Opnd, S),
{C#cache.expr, count_size(C#cache.size, S1)}.
visit_1(Opnd, S) ->
case st__lookup_opnd_cache(Opnd#opnd.loc, S) of
error ->
%% Use a new, passive, size counter for visiting operands,
%% and use the effort counter of the context of the operand.
%% It turns out that if the latter is active, it must be the
%% same object as the one currently used, and if it is
%% passive, it does not matter if it is the same object as
%% any other counter.
Effort = Opnd#opnd.effort,
Active = counter__is_active(Effort),
S1 = case Active of
true ->
S; % don't change effort counter
false ->
st__set_effort(Effort, S)
end,
S2 = new_passive_size(get_size_limit(S1), S1),
%% Visit the expression and cache the result, along with the
%% final value of the size counter.
{E, S3} = i(Opnd#opnd.expr, value, Opnd#opnd.ren,
Opnd#opnd.env, S2),
Size = get_size_value(S3),
C = #cache{expr = E, size = Size},
S4 = revert_size(S, st__set_opnd_cache(Opnd#opnd.loc, C,
S3)),
case Active of
true ->
{C, S4}; % keep using the same effort counter
false ->
{C, revert_effort(S, S4)}
end;
{ok, C} ->
{C, S}
end.
%% Create a pattern matching template for an expression. A template
%% contains only data constructors (including atomic ones) and
%% variables, and compound literals are not folded into a single node.
%% Each node in the template is annotated with the variable which holds
%% the corresponding subexpression; these are new, unique variables not
%% existing in the given `Env'. Returns `{Template, Variables, NewEnv}',
%% where `Variables' is the list of all variables corresponding to nodes
%% in the template *listed in reverse dependency order*, and `NewEnv' is
%% `Env' augmented with mappings from the variable names to
%% subexpressions of `E' (not #ref{} structures!) rewritten so that no
%% computations are duplicated. `Variables' is guaranteed to be nonempty
%% - at least the root node will always be bound to a new variable.
make_template(E, Env) ->
make_template(E, [], Env).
make_template(E, Vs0, Env0) ->
case is_data(E) of
true ->
{Ts, {Vs1, Env1}} = mapfoldl(
fun (E, {Vs0, Env0}) ->
{T, Vs1, Env1} =
make_template(E, Vs0,
Env0),
{T, {Vs1, Env1}}
end,
{Vs0, Env0}, data_es(E)),
T = make_data_skel(data_type(E), Ts),
E1 = update_data(E, data_type(E),
[hd(get_ann(T)) || T <- Ts]),
V = new_var(Env1),
Env2 = env__bind(var_name(V), E1, Env1),
{set_ann(T, [V]), [V | Vs1], Env2};
false ->
case type(E) of
seq ->
%% For a sequencing, we can rebind the variable used
%% for the body, and pass on the template as it is.
{T, Vs1, Env1} = make_template(seq_body(E), Vs0,
Env0),
V = var_name(hd(get_ann(T))),
E1 = update_c_seq(E, seq_arg(E), env__get(V, Env1)),
Env2 = env__bind(V, E1, Env1),
{T, Vs1, Env2};
_ ->
V = new_var(Env0),
Env1 = env__bind(var_name(V), E, Env0),
{set_ann(V, [V]), [V | Vs0], Env1}
end
end.
%% Two clauses are equivalent if their bodies are equivalent expressions
%% given that the respective pattern variables are local.
equivalent_clauses([]) ->
true;
equivalent_clauses([C | Cs]) ->
Env = cerl_trees:variables(c_values(clause_pats(C))),
equivalent_clauses_1(clause_body(C), Cs, Env).
equivalent_clauses_1(E, [C | Cs], Env) ->
Env1 = cerl_trees:variables(c_values(clause_pats(C))),
case equivalent(E, clause_body(C), ordsets:union(Env, Env1)) of
true ->
equivalent_clauses_1(E, Cs, Env);
false ->
false
end;
equivalent_clauses_1(_, [], _Env) ->
true.
%% Two expressions are equivalent if and only if they yield the same
%% value and has the same side effects in the same order. Currently, we
%% only accept equality between constructors (constants) and nonlocal
%% variables, since this should cover most cases of interest. If a
%% variable is locally bound in one expression, it cannot be equivalent
%% to one with the same name in the other expression, so we need not
%% keep track of two environments.
equivalent(E1, E2, Env) ->
case is_data(E1) of
true ->
case is_data(E2) of
true ->
T1 = {data_type(E1), data_arity(E1)},
T2 = {data_type(E2), data_arity(E2)},
%% Note that we must test for exact equality.
if T1 =:= T2 ->
equivalent_lists(data_es(E1), data_es(E2),
Env);
true ->
false
end;
false ->
false
end;
false ->
case type(E1) of
var ->
case is_c_var(E2) of
true ->
N1 = var_name(E1),
N2 = var_name(E2),
if N1 =:= N2 ->
not ordsets:is_element(N1, Env);
true ->
false
end;
false ->
false
end;
_ ->
%% Other constructs are not being considered.
false
end
end.
equivalent_lists([E1 | Es1], [E2 | Es2], Env) ->
equivalent(E1, E2, Env) and equivalent_lists(Es1, Es2, Env);
equivalent_lists([], [], _) ->
true;
equivalent_lists(_, _, _) ->
false.
%% Return `false' or `{true, EffectExpr, ValueExpr}'. The environment is
%% passed for new-variable generation.
reduce_bif_call(M, F, As, Env) ->
reduce_bif_call_1(M, F, length(As), As, Env).
reduce_bif_call_1(erlang, element, 2, [X, Y], _Env) ->
case is_c_int(X) and is_c_tuple(Y) of
true ->
%% We are free to change the relative evaluation order of
%% the elements, so lifting out a particular element is OK.
T = list_to_tuple(tuple_es(Y)),
N = int_val(X),
if integer(N), N > 0, N =< size(T) ->
E = element(N, T),
Es = tuple_to_list(setelement(N, T, void())),
{true, make_seq(c_tuple(Es), E)};
true ->
false
end;
false ->
false
end;
reduce_bif_call_1(erlang, hd, 1, [X], _Env) ->
case is_c_cons(X) of
true ->
%% Cf. `element/2' above.
{true, make_seq(cons_tl(X), cons_hd(X))};
false ->
false
end;
reduce_bif_call_1(erlang, length, 1, [X], _Env) ->
case is_c_list(X) of
true ->
%% Cf. `erlang:size/1' below.
{true, make_seq(X, c_int(list_length(X)))};
false ->
false
end;
reduce_bif_call_1(erlang, list_to_tuple, 1, [X], _Env) ->
case is_c_list(X) of
true ->
%% This does not actually preserve all the evaluation order
%% constraints of the list, but I don't imagine that it will
%% be a problem.
{true, c_tuple(list_elements(X))};
false ->
false
end;
reduce_bif_call_1(erlang, setelement, 3, [X, Y, Z], Env) ->
case is_c_int(X) and is_c_tuple(Y) of
true ->
%% Here, unless `Z' is a simple expression, we must bind it
%% to a new variable, because in that case, `Z' must be
%% evaluated before any part of `Y'.
T = list_to_tuple(tuple_es(Y)),
N = int_val(X),
if integer(N), N > 0, N =< size(T) ->
E = element(N, T),
case is_simple(Z) of
true ->
Es = tuple_to_list(setelement(N, T, Z)),
{true, make_seq(E, c_tuple(Es))};
false ->
V = new_var(Env),
Es = tuple_to_list(setelement(N, T, V)),
E1 = make_seq(E, c_tuple(Es)),
{true, c_let([V], Z, E1)}
end;
true ->
false
end;
false ->
false
end;
reduce_bif_call_1(erlang, size, 1, [X], _Env) ->
case is_c_tuple(X) of
true ->
%% Just evaluate the tuple for effect and use the size (the
%% arity) as the result.
{true, make_seq(X, c_int(tuple_arity(X)))};
false ->
false
end;
reduce_bif_call_1(erlang, tl, 1, [X], _Env) ->
case is_c_cons(X) of
true ->
%% Cf. `element/2' above.
{true, make_seq(cons_hd(X), cons_tl(X))};
false ->
false
end;
reduce_bif_call_1(erlang, tuple_to_list, 1, [X], _Env) ->
case is_c_tuple(X) of
true ->
%% This actually introduces slightly stronger constraints on
%% the evaluation order of the subexpressions.
{true, make_list(tuple_es(X))};
false ->
false
end;
reduce_bif_call_1(_M, _F, _A, _As, _Env) ->
false.
effort_is_active(S) ->
counter__is_active(st__get_effort(S)).
size_is_active(S) ->
counter__is_active(st__get_size(S)).
get_effort_limit(S) ->
counter__limit(st__get_effort(S)).
new_active_effort(Limit, S) ->
st__set_effort(counter__new_active(Limit), S).
revert_effort(S1, S2) ->
st__set_effort(st__get_effort(S1), S2).
new_active_size(Limit, S) ->
st__set_size(counter__new_active(Limit), S).
new_passive_size(Limit, S) ->
st__set_size(counter__new_passive(Limit), S).
revert_size(S1, S2) ->
st__set_size(st__get_size(S1), S2).
count_effort(N, S) ->
C = st__get_effort(S),
C1 = counter__add(N, C, effort, S),
case debug_counters() of
true ->
case counter__is_active(C1) of
true ->
V = counter__value(C1),
case V > get(counter_effort_max) of
true ->
put(counter_effort_max, V);
false ->
ok
end;
false ->
ok
end;
_ ->
ok
end,
st__set_effort(C1, S).
count_size(N, S) ->
C = st__get_size(S),
C1 = counter__add(N, C, size, S),
case debug_counters() of
true ->
case counter__is_active(C1) of
true ->
V = counter__value(C1),
case V > get(counter_size_max) of
true ->
put(counter_size_max, V);
false ->
ok
end;
false ->
ok
end;
_ ->
ok
end,
st__set_size(C1, S).
get_size_value(S) ->
counter__value(st__get_size(S)).
get_size_limit(S) ->
counter__limit(st__get_size(S)).
kill_id_anns([{'id',_} | As]) ->
kill_id_anns(As);
kill_id_anns([A | As]) ->
[A | kill_id_anns(As)];
kill_id_anns([]) ->
[].
%% =====================================================================
%% General utilities
max(X, Y) when X > Y -> X;
max(_, Y) -> Y.
%% The atom `ok', is widely used in Erlang for "void" values.
void() -> abstract(ok).
is_simple(E) ->
case type(E) of
literal -> true;
var -> true;
'fun' -> true;
_ -> false
end.
get_components(N, E) ->
case type(E) of
values ->
Es = values_es(E),
if length(Es) == N ->
{true, Es};
true ->
false
end;
_ when N == 1 ->
{true, [E]};
_ ->
false
end.
all_static([E | Es]) ->
case is_literal(result(E)) of
true ->
all_static(Es);
false ->
false
end;
all_static([]) ->
true.
set_clause_bodies([C | Cs], B) ->
[update_c_clause(C, clause_pats(C), clause_guard(C), B)
| set_clause_bodies(Cs, B)];
set_clause_bodies([], _) ->
[].
filename([C | T]) when integer(C), C > 0, C =< 255 ->
[C | filename(T)];
filename([H|T]) ->
filename(H) ++ filename(T);
filename([]) ->
[];
filename(N) when atom(N) ->
atom_to_list(N);
filename(N) ->
report_error("bad filename: `~P'.", [N, 25]),
exit(error).
%% =====================================================================
%% Abstract datatype: renaming()
ren__identity() ->
dict:new().
ren__add(X, Y, Ren) ->
dict:store(X, Y, Ren).
ren__map(X, Ren) ->
case dict:find(X, Ren) of
{ok, Y} ->
Y;
error ->
X
end.
ren__add_identity(X, Ren) ->
dict:erase(X, Ren).
%% =====================================================================
%% Abstract datatype: environment()
env__empty() ->
rec_env:empty().
env__bind(Key, Val, Env) ->
rec_env:bind(Key, Val, Env).
%% `Es' should have type `[{Key, Val}]', and `Fun' should have type
%% `(Val, Env) -> T', mapping a value together with the recursive
%% environment itself to some term `T' to be returned when the entry is
%% looked up.
env__bind_recursive(Ks, Vs, F, Env) ->
rec_env:bind_recursive(Ks, Vs, F, Env).
env__lookup(Key, Env) ->
rec_env:lookup(Key, Env).
env__get(Key, Env) ->
rec_env:get(Key, Env).
env__is_defined(Key, Env) ->
rec_env:is_defined(Key, Env).
env__new_vname(Env) ->
rec_env:new_key(Env).
env__new_fname(A, N, Env) ->
rec_env:new_key(fun (X) ->
S = integer_to_list(X),
{list_to_atom(atom_to_list(A) ++ "_" ++ S),
N}
end, Env).
%% =====================================================================
%% Abstract datatype: state()
-record(state, {free, % next free location
size, % size counter
effort, % effort counter
cache, % operand expression cache
var_flags, % flags for variables (#ref-structures)
opnd_flags, % flags for operands
app_flags}). % flags for #app-structures
%% Note that we do not have a `var_assigned' flag, since there is no
%% destructive assignment in Erlang. In the original algorithm, the
%% "residual-referenced"-flags of the previous inlining pass (or
%% initialization pass) are used as the "source-referenced"-flags for
%% the subsequent pass. The latter may then be used as a safe
%% approximation whenever we need to base a decision on whether or not a
%% particular variable or function variable could be referenced in the
%% program being generated, and computation of the new
%% "residual-referenced" flag for that variable is not yet finished. In
%% the present algorithm, this can only happen in the presence of
%% variable assignments, which do not exist in Erlang. Therefore, we do
%% not keep "source-referenced" flags for residual-code references in
%% our implementation.
%%
%% The "inner-pending" flag tells us whether we are already in the
%% process of visiting a particular operand, and the "outer-pending"
%% flag whether we are in the process of inlining a propagated
%% functional value. The "pending flags" are really counters limiting
%% the number of times an operand may be inlined recursively, causing
%% loop unrolling; however, unrolling more than one iteration does not
%% work offhand in the present implementation. (TODO: find out why.)
%% Note that the initial value must be greater than zero in order for
%% any inlining at all to be done.
%% Flags are stored in ETS-tables, one table for each class. The second
%% element in each stored tuple is the key (the "label").
-record(var_flags, {lab, referenced = false}).
-record(opnd_flags, {lab, inner_pending = 1, outer_pending = 1,
effect = false}).
-record(app_flags, {lab, inlined = false}).
st__new(Effort, Size) ->
#state{free = 0,
size = counter__new_passive(Size),
effort = counter__new_passive(Effort),
cache = dict:new(),
var_flags = ets:new(var, [set, private, {keypos, 2}]),
opnd_flags = ets:new(opnd, [set, private, {keypos, 2}]),
app_flags = ets:new(app, [set, private, {keypos, 2}])}.
st__new_loc(S) ->
N = S#state.free,
{N, S#state{free = N + 1}}.
st__get_effort(S) ->
S#state.effort.
st__set_effort(C, S) ->
S#state{effort = C}.
st__get_size(S) ->
S#state.size.
st__set_size(C, S) ->
S#state{size = C}.
st__set_var_referenced(L, S) ->
T = S#state.var_flags,
[F] = ets:lookup(T, L),
ets:insert(T, F#var_flags{referenced = true}),
S.
st__get_var_referenced(L, S) ->
ets:lookup_element(S#state.var_flags, L, #var_flags.referenced).
st__lookup_opnd_cache(L, S) ->
dict:find(L, S#state.cache).
%% Note that setting the cache should only be done once.
st__set_opnd_cache(L, C, S) ->
S#state{cache = dict:store(L, C, S#state.cache)}.
st__set_opnd_effect(L, S) ->
T = S#state.opnd_flags,
[F] = ets:lookup(T, L),
ets:insert(T, F#opnd_flags{effect = true}),
S.
st__get_opnd_effect(L, S) ->
ets:lookup_element(S#state.opnd_flags, L, #opnd_flags.effect).
st__set_app_inlined(L, S) ->
T = S#state.app_flags,
[F] = ets:lookup(T, L),
ets:insert(T, F#app_flags{inlined = true}),
S.
st__clear_app_inlined(L, S) ->
T = S#state.app_flags,
[F] = ets:lookup(T, L),
ets:insert(T, F#app_flags{inlined = false}),
S.
st__get_app_inlined(L, S) ->
ets:lookup_element(S#state.app_flags, L, #app_flags.inlined).
%% The pending-flags are initialized by `st__new_opnd_loc' below.
st__test_inner_pending(L, S) ->
T = S#state.opnd_flags,
P = ets:lookup_element(T, L, #opnd_flags.inner_pending),
P =< 0.
st__mark_inner_pending(L, S) ->
ets:update_counter(S#state.opnd_flags, L,
{#opnd_flags.inner_pending, -1}),
S.
st__clear_inner_pending(L, S) ->
ets:update_counter(S#state.opnd_flags, L,
{#opnd_flags.inner_pending, 1}),
S.
st__test_outer_pending(L, S) ->
T = S#state.opnd_flags,
P = ets:lookup_element(T, L, #opnd_flags.outer_pending),
P =< 0.
st__mark_outer_pending(L, S) ->
ets:update_counter(S#state.opnd_flags, L,
{#opnd_flags.outer_pending, -1}),
S.
st__clear_outer_pending(L, S) ->
ets:update_counter(S#state.opnd_flags, L,
{#opnd_flags.outer_pending, 1}),
S.
st__new_app_loc(S) ->
V = {L, _S1} = st__new_loc(S),
ets:insert(S#state.app_flags, #app_flags{lab = L}),
V.
st__new_ref_loc(S) ->
V = {L, _S1} = st__new_loc(S),
ets:insert(S#state.var_flags, #var_flags{lab = L}),
V.
st__new_opnd_loc(S) ->
V = {L, _S1} = st__new_loc(S),
ets:insert(S#state.opnd_flags, #opnd_flags{lab = L}),
V.
%% =====================================================================
%% Abstract datatype: counter()
%%
%% `counter__add' throws `{counter_exceeded, Type, Data}' if the
%% resulting counter value would exceed the limit for the counter in
%% question (`Type' and `Data' are given by the user).
-record(counter, {active, value, limit}).
counter__new_passive(Limit) when Limit > 0 ->
{0, Limit}.
counter__new_active(Limit) when Limit > 0 ->
{Limit, Limit}.
%% Active counters have values > 0 internally; passive counters start at
%% zero. The 'limit' field is only accessed by the 'counter__limit'
%% function.
counter__is_active({C, _}) ->
C > 0.
counter__limit({_, L}) ->
L.
counter__value({N, L}) ->
if N > 0 ->
L - N;
true ->
-N
end.
counter__add(N, {V, L}, Type, Data) ->
N1 = V - N,
if V > 0, N1 =< 0 ->
case debug_counters() of
true ->
case Type of
effort ->
put(counter_effort_triggers,
get(counter_effort_triggers) + 1);
size ->
put(counter_size_triggers,
get(counter_size_triggers) + 1)
end;
_ ->
ok
end,
throw({counter_exceeded, Type, Data});
true ->
{N1, L}
end.
%% =====================================================================
%% Reporting
% report_internal_error(S) ->
% report_internal_error(S, []).
report_internal_error(S, Vs) ->
report_error("internal error: " ++ S, Vs).
report_error(D) ->
report_error(D, []).
report_error({F, L, D}, Vs) ->
report({F, L, {error, D}}, Vs);
report_error(D, Vs) ->
report({error, D}, Vs).
report_warning(D) ->
report_warning(D, []).
report_warning({F, L, D}, Vs) ->
report({F, L, {warning, D}}, Vs);
report_warning(D, Vs) ->
report({warning, D}, Vs).
report(D, Vs) ->
io:put_chars(format(D, Vs)).
format({error, D}, Vs) ->
["error: ", format(D, Vs)];
format({warning, D}, Vs) ->
["warning: ", format(D, Vs)];
format({"", L, D}, Vs) when integer(L), L > 0 ->
[io_lib:fwrite("~w: ", [L]), format(D, Vs)];
format({"", _L, D}, Vs) ->
format(D, Vs);
format({F, L, D}, Vs) when integer(L), L > 0 ->
[io_lib:fwrite("~s:~w: ", [filename(F), L]), format(D, Vs)];
format({F, _L, D}, Vs) ->
[io_lib:fwrite("~s: ", [filename(F)]), format(D, Vs)];
format(S, Vs) when list(S) ->
[io_lib:fwrite(S, Vs), $\n].
%% =====================================================================
