= tagZero xs #-}
So tagZero's RHS mentions pmap, and pmap's RULE mentions tagZero.
However, tagZero can only be inlined in phase 1 and later, while
the RULE is only active *before* phase 1. So there's no problem.
To make this work, we look for the RHS free vars only for
*active* rules. That's the reason for the occ_rule_act field
of the OccEnv.
* Note [Weak loop breakers]
~~~~~~~~~~~~~~~~~~~~~~~~~
There is a last nasty wrinkle. Suppose we have
Rec { f = f_rhs
RULE f [] = g
h = h_rhs
g = h
...more...
}
Remember that we simplify the RULES before any RHS (see Note
[Rules are visible in their own rec group] above).
So we must *not* postInlineUnconditionally 'g', even though
its RHS turns out to be trivial. (I'm assuming that 'g' is
not choosen as a loop breaker.) Why not? Because then we
drop the binding for 'g', which leaves it out of scope in the
RULE!
Here's a somewhat different example of the same thing
Rec { g = h
; h = ...f...
; f = f_rhs
RULE f [] = g }
Here the RULE is "below" g, but we *still* can't postInlineUnconditionally
g, because the RULE for f is active throughout. So the RHS of h
might rewrite to h = ...g...
So g must remain in scope in the output program!
We "solve" this by:
Make g a "weak" loop breaker (OccInfo = IAmLoopBreaker True)
iff g is a "missing free variable" of the Rec group
A "missing free variable" x is one that is mentioned in an RHS or
INLINE or RULE of a binding in the Rec group, but where the
dependency on x may not show up in the loop_breaker_edges (see
note [Choosing loop breakers} above).
A normal "strong" loop breaker has IAmLoopBreaker False. So
Inline postInlineUnconditionally
IAmLoopBreaker False no no
IAmLoopBreaker True yes no
other yes yes
The **sole** reason for this kind of loop breaker is so that
postInlineUnconditionally does not fire. Ugh. (Typically it'll
inline via the usual callSiteInline stuff, so it'll be dead in the
next pass, so the main Ugh is the tiresome complication.)
Note [Rules for imported functions]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider this
f = /\a. B.g a
RULE B.g Int = 1 + f Int
Note that
* The RULE is for an imported function.
* f is non-recursive
Now we
can get
f Int --> B.g Int Inlining f
--> 1 + f Int Firing RULE
and so the simplifier goes into an infinite loop. This
would not happen if the RULE was for a local function,
because we keep track of dependencies through rules. But
that is pretty much impossible to do for imported Ids. Suppose
f's definition had been
f = /\a. C.h a
where (by some long and devious process), C.h eventually inlines to
B.g. We could only spot such loops by exhaustively following
unfoldings of C.h etc, in case we reach B.g, and hence (via the RULE)
f.
Note that RULES for imported functions are important in practice; they
occur a lot in the libraries.
We regard this potential infinite loop as a *programmer* error.
It's up the programmer not to write silly rules like
RULE f x = f x
and the example above is just a more complicated version.
Note [Preventing loops due to imported functions rules]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider:
import GHC.Base (foldr)
{-# RULES "filterList" forall p. foldr (filterFB (:) p) [] = filter p #-}
filter p xs = build (\c n -> foldr (filterFB c p) n xs)
filterFB c p = ...
f = filter p xs
Note that filter is not a loop-breaker, so what happens is:
f = filter p xs
= {inline} build (\c n -> foldr (filterFB c p) n xs)
= {inline} foldr (filterFB (:) p) [] xs
= {RULE} filter p xs
We are in an infinite loop.
A more elaborate example (that I actually saw in practice when I went to
mark GHC.List.filter as INLINABLE) is as follows. Say I have this module:
{-# LANGUAGE RankNTypes #-}
module GHCList where
import Prelude hiding (filter)
import GHC.Base (build)
{-# INLINABLE filter #-}
filter :: (a -> Bool) -> [a] -> [a]
filter p [] = []
filter p (x:xs) = if p x then x : filter p xs else filter p xs
{-# NOINLINE [0] filterFB #-}
filterFB :: (a -> b -> b) -> (a -> Bool) -> a -> b -> b
filterFB c p x r | p x = x `c` r
| otherwise = r
{-# RULES
"filter" [~1] forall p xs. filter p xs = build (\c n -> foldr
(filterFB c p) n xs)
"filterList" [1] forall p. foldr (filterFB (:) p) [] = filter p
#-}
Then (because RULES are applied inside INLINABLE unfoldings, but inlinings
are not), the unfolding given to "filter" in the interface file will be:
filter p [] = []
filter p (x:xs) = if p x then x : build (\c n -> foldr (filterFB c p) n xs)
else build (\c n -> foldr (filterFB c p) n xs
Note that because this unfolding does not mention "filter", filter is not
marked as a strong loop breaker. Therefore at a use site in another module:
filter p xs
= {inline}
case xs of [] -> []
(x:xs) -> if p x then x : build (\c n -> foldr (filterFB c p) n xs)
else build (\c n -> foldr (filterFB c p) n xs)
build (\c n -> foldr (filterFB c p) n xs)
= {inline} foldr (filterFB (:) p) [] xs
= {RULE} filter p xs
And we are in an infinite loop again, except that this time the loop is producing an
infinitely large *term* (an unrolling of filter) and so the simplifier finally
dies with "ticks exhausted"
Because of this problem, we make a small change in the occurrence analyser
designed to mark functions like "filter" as strong loop breakers on the basis that:
1. The RHS of filter mentions the local function "filterFB"
2. We have a rule which mentions "filterFB" on the LHS and "filter" on the RHS
So for each RULE for an *imported* function we are going to add
dependency edges between the *local* FVS of the rule LHS and the
*local* FVS of the rule RHS. We don't do anything special for RULES on
local functions because the standard occurrence analysis stuff is
pretty good at getting loop-breakerness correct there.
It is important to note that even with this extra hack we aren't always going to get
things right. For example, it might be that the rule LHS mentions an imported Id,
and another module has a RULE that can rewrite that imported Id to one of our local
Ids.
Note [Specialising imported functions]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
BUT for *automatically-generated* rules, the programmer can't be
responsible for the "programmer error" in Note [Rules for imported
functions]. In paricular, consider specialising a recursive function
defined in another module. If we specialise a recursive function B.g,
we get
g_spec = .....(B.g Int).....
RULE B.g Int = g_spec
Here, g_spec doesn't look recursive, but when the rule fires, it
becomes so. And if B.g was mutually recursive, the loop might
not be as obvious as it is here.
To avoid this,
* When specialising a function that is a loop breaker,
give a NOINLINE pragma to the specialised function
Note [Glomming]
~~~~~~~~~~~~~~~
RULES for imported Ids can make something at the top refer to something at the bottom:
f = \x -> B.g (q x)
h = \y -> 3
RULE: B.g (q x) = h x
Applying this rule makes f refer to h, although f doesn't appear to
depend on h. (And, as in Note [Rules for imported functions], the
dependency might be more indirect. For example, f might mention C.t
rather than B.g, where C.t eventually inlines to B.g.)
NOTICE that this cannot happen for rules whose head is a
locally-defined function, because we accurately track dependencies
through RULES. It only happens for rules whose head is an imported
function (B.g in the example above).
Solution:
- When simplifying, bring all top level identifiers into
scope at the start, ignoring the Rec/NonRec structure, so
that when 'h' pops up in f's rhs, we find it in the in-scope set
(as the simplifier generally expects). This happens in simplTopBinds.
- In the occurrence analyser, if there are any out-of-scope
occurrences that pop out of the top, which will happen after
firing the rule: f = \x -> h x
h = \y -> 3
then just glom all the bindings into a single Rec, so that
the *next* iteration of the occurrence analyser will sort
them all out. This part happens in occurAnalysePgm.
------------------------------------------------------------
Note [Inline rules]
~~~~~~~~~~~~~~~~~~~
None of the above stuff about RULES applies to Inline Rules,
stored in a CoreUnfolding. The unfolding, if any, is simplified
at the same time as the regular RHS of the function (ie *not* like
Note [Rules are visible in their own rec group]), so it should be
treated *exactly* like an extra RHS.
Or, rather, when computing loop-breaker edges,
* If f has an INLINE pragma, and it is active, we treat the
INLINE rhs as f's rhs
* If it's inactive, we treat f as having no rhs
* If it has no INLINE pragma, we look at f's actual rhs
There is a danger that we'll be sub-optimal if we see this
f = ...f...
[INLINE f = ..no f...]
where f is recursive, but the INLINE is not. This can just about
happen with a sufficiently odd set of rules; eg
foo :: Int -> Int
{-# INLINE [1] foo #-}
foo x = x+1
bar :: Int -> Int
{-# INLINE [1] bar #-}
bar x = foo x + 1
{-# RULES "foo" [~1] forall x. foo x = bar x #-}
Here the RULE makes bar recursive; but it's INLINE pragma remains
non-recursive. It's tempting to then say that 'bar' should not be
a loop breaker, but an attempt to do so goes wrong in two ways:
a) We may get
$df = ...$cfoo...
$cfoo = ...$df....
[INLINE $cfoo = ...no-$df...]
But we want $cfoo to depend on $df explicitly so that we
put the bindings in the right order to inline $df in $cfoo
and perhaps break the loop altogether. (Maybe this
b)
Example [eftInt]
~~~~~~~~~~~~~~~
Example (from GHC.Enum):
eftInt :: Int# -> Int# -> [Int]
eftInt x y = ...(non-recursive)...
{-# INLINE [0] eftIntFB #-}
eftIntFB :: (Int -> r -> r) -> r -> Int# -> Int# -> r
eftIntFB c n x y = ...(non-recursive)...
{-# RULES
"eftInt" [~1] forall x y. eftInt x y = build (\ c n -> eftIntFB c n x y)
"eftIntList" [1] eftIntFB (:) [] = eftInt
#-}
Note [Specialisation rules]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider this group, which is typical of what SpecConstr builds:
fs a = ....f (C a)....
f x = ....f (C a)....
{-# RULE f (C a) = fs a #-}
So 'f' and 'fs' are in the same Rec group (since f refers to fs via its RULE).
But watch out! If 'fs' is not chosen as a loop breaker, we may get an infinite loop:
- the RULE is applied in f's RHS (see Note [Self-recursive rules] in Simplify
- fs is inlined (say it's small)
- now there's another opportunity to apply the RULE
This showed up when compiling Control.Concurrent.Chan.getChanContents.
\begin{code}
type Node details = (details, Unique, [Unique])
data Details
= ND { nd_bndr :: Id
, nd_rhs :: CoreExpr
, nd_uds :: UsageDetails
, nd_inl :: IdSet
, nd_weak :: IdSet
, nd_active_rule_fvs :: IdSet
}
instance Outputable Details where
ppr nd = ptext (sLit "ND") <> braces
(sep [ ptext (sLit "bndr =") <+> ppr (nd_bndr nd)
, ptext (sLit "uds =") <+> ppr (nd_uds nd)
, ptext (sLit "inl =") <+> ppr (nd_inl nd)
, ptext (sLit "weak =") <+> ppr (nd_weak nd)
, ptext (sLit "rule =") <+> ppr (nd_active_rule_fvs nd)
])
makeNode :: OccEnv -> IdEnv IdSet -> VarSet -> (Var, CoreExpr) -> Node Details
makeNode env imp_rules_edges bndr_set (bndr, rhs)
= (details, varUnique bndr, keysUFM node_fvs)
where
details = ND { nd_bndr = bndr
, nd_rhs = rhs'
, nd_uds = rhs_usage3
, nd_weak = node_fvs `minusVarSet` inl_fvs
, nd_inl = inl_fvs
, nd_active_rule_fvs = active_rule_fvs }
(rhs_usage1, rhs') = occAnalRecRhs env rhs
rhs_usage2 = addIdOccs rhs_usage1 all_rule_fvs
rhs_usage3 = case mb_unf_fvs of
Just unf_fvs -> addIdOccs rhs_usage2 unf_fvs
Nothing -> rhs_usage2
node_fvs = udFreeVars bndr_set rhs_usage3
is_active = occ_rule_act env :: Activation -> Bool
rules = filterOut isBuiltinRule (idCoreRules bndr)
rules_w_fvs :: [(Activation, VarSet)]
rules_w_fvs = maybe id (\ids -> ((AlwaysActive, ids):)) (lookupVarEnv imp_rules_edges bndr)
[ (ru_act rule, fvs)
| rule <- rules
, let fvs = exprFreeVars (ru_rhs rule)
`delVarSetList` ru_bndrs rule
, not (isEmptyVarSet fvs) ]
all_rule_fvs = foldr (unionVarSet . snd) rule_lhs_fvs rules_w_fvs
rule_lhs_fvs = foldr (unionVarSet . (\ru -> exprsFreeVars (ru_args ru)
`delVarSetList` ru_bndrs ru))
emptyVarSet rules
active_rule_fvs = unionVarSets [fvs | (a,fvs) <- rules_w_fvs, is_active a]
unf = realIdUnfolding bndr
mb_unf_fvs = stableUnfoldingVars unf
inl_fvs = case mb_unf_fvs of
Nothing -> udFreeVars bndr_set rhs_usage1
Just unf_fvs -> unf_fvs
occAnalRec :: SCC (Node Details)
-> (UsageDetails, [CoreBind])
-> (UsageDetails, [CoreBind])
occAnalRec (AcyclicSCC (ND { nd_bndr = bndr, nd_rhs = rhs, nd_uds = rhs_uds}, _, _))
(body_uds, binds)
| not (bndr `usedIn` body_uds)
= (body_uds, binds)
| otherwise
= (body_uds' +++ rhs_uds,
NonRec tagged_bndr rhs : binds)
where
(body_uds', tagged_bndr) = tagBinder body_uds bndr
occAnalRec (CyclicSCC nodes) (body_uds, binds)
| not (any (`usedIn` body_uds) bndrs)
= (body_uds, binds)
| otherwise
=
(final_uds, Rec pairs : binds)
where
bndrs = [b | (ND { nd_bndr = b }, _, _) <- nodes]
bndr_set = mkVarSet bndrs
tagged_nodes = map tag_node nodes
total_uds = foldl add_uds body_uds nodes
final_uds = total_uds `minusVarEnv` bndr_set
add_uds usage_so_far (nd, _, _) = usage_so_far +++ nd_uds nd
tag_node :: Node Details -> Node Details
tag_node (details@ND { nd_bndr = bndr }, k, ks)
= (details { nd_bndr = setBinderOcc total_uds bndr }, k, ks)
pairs :: [(Id,CoreExpr)]
pairs | isEmptyVarSet weak_fvs = reOrderNodes 0 bndr_set weak_fvs tagged_nodes []
| otherwise = loopBreakNodes 0 bndr_set weak_fvs loop_breaker_edges []
weak_fvs :: VarSet
weak_fvs = foldr (unionVarSet . nd_weak . fstOf3) emptyVarSet nodes
loop_breaker_edges = map mk_node tagged_nodes
mk_node (details@(ND { nd_inl = inl_fvs }), k, _)
= (details, k, keysUFM (extendFvs_ rule_fv_env inl_fvs))
rule_fv_env :: IdEnv IdSet
rule_fv_env = transClosureFV (mkVarEnv init_rule_fvs)
init_rule_fvs
= [ (b, trimmed_rule_fvs)
| (ND { nd_bndr = b, nd_active_rule_fvs = rule_fvs },_,_) <- nodes
, let trimmed_rule_fvs = rule_fvs `intersectVarSet` bndr_set
, not (isEmptyVarSet trimmed_rule_fvs)]
\end{code}
@loopBreakSCC@ is applied to the list of (binder,rhs) pairs for a cyclic
strongly connected component (there's guaranteed to be a cycle). It returns the
same pairs, but
a) in a better order,
b) with some of the Ids having a IAmALoopBreaker pragma
The "loop-breaker" Ids are sufficient to break all cycles in the SCC. This means
that the simplifier can guarantee not to loop provided it never records an inlining
for these no-inline guys.
Furthermore, the order of the binds is such that if we neglect dependencies
on the no-inline Ids then the binds are topologically sorted. This means
that the simplifier will generally do a good job if it works from top bottom,
recording inlinings for any Ids which aren't marked as "no-inline" as it goes.
\begin{code}
type Binding = (Id,CoreExpr)
mk_loop_breaker :: Node Details -> Binding
mk_loop_breaker (ND { nd_bndr = bndr, nd_rhs = rhs}, _, _)
= (setIdOccInfo bndr strongLoopBreaker, rhs)
mk_non_loop_breaker :: VarSet -> Node Details -> Binding
mk_non_loop_breaker used_in_rules (ND { nd_bndr = bndr, nd_rhs = rhs}, _, _)
| bndr `elemVarSet` used_in_rules = (setIdOccInfo bndr weakLoopBreaker, rhs)
| otherwise = (bndr, rhs)
udFreeVars :: VarSet -> UsageDetails -> VarSet
udFreeVars bndrs uds = intersectUFM_C (\b _ -> b) bndrs uds
loopBreakNodes :: Int
-> VarSet
-> VarSet
-> [Node Details]
-> [Binding]
-> [Binding]
loopBreakNodes depth bndr_set weak_fvs nodes binds
= go (stronglyConnCompFromEdgedVerticesR nodes) binds
where
go [] binds = binds
go (scc:sccs) binds = loop_break_scc scc (go sccs binds)
loop_break_scc scc binds
= case scc of
AcyclicSCC node -> mk_non_loop_breaker weak_fvs node : binds
CyclicSCC [node] -> mk_loop_breaker node : binds
CyclicSCC nodes -> reOrderNodes depth bndr_set weak_fvs nodes binds
reOrderNodes :: Int -> VarSet -> VarSet -> [Node Details] -> [Binding] -> [Binding]
reOrderNodes _ _ _ [] _ = panic "reOrderNodes"
reOrderNodes depth bndr_set weak_fvs (node : nodes) binds
=
loopBreakNodes new_depth bndr_set weak_fvs unchosen $
(map mk_loop_breaker chosen_nodes ++ binds)
where
(chosen_nodes, unchosen) = choose_loop_breaker (score node) [node] [] nodes
approximate_loop_breaker = depth >= 2
new_depth | approximate_loop_breaker = 0
| otherwise = depth+1
choose_loop_breaker :: Int
-> [Node Details]
-> [Node Details]
-> [Node Details]
-> ([Node Details], [Node Details])
choose_loop_breaker _ loop_nodes acc []
= (loop_nodes, acc)
choose_loop_breaker loop_sc loop_nodes acc (node : nodes)
| sc < loop_sc
= choose_loop_breaker sc [node] (loop_nodes ++ acc) nodes
| approximate_loop_breaker && sc == loop_sc
= choose_loop_breaker loop_sc (node : loop_nodes) acc nodes
| otherwise
= choose_loop_breaker loop_sc loop_nodes (node : acc) nodes
where
sc = score node
score :: Node Details -> Int
score (ND { nd_bndr = bndr, nd_rhs = rhs }, _, _)
| not (isId bndr) = 100
| isDFunId bndr = 9
| Just _ <- isStableCoreUnfolding_maybe (idUnfolding bndr)
= 3
| is_con_app rhs = 5
| exprIsTrivial rhs = 10
| isOneOcc (idOccInfo bndr) = 2
| canUnfold (realIdUnfolding bndr) = 1
| otherwise = 0
is_con_app (Var v) = isConLikeId v
is_con_app (App f _) = is_con_app f
is_con_app (Lam _ e) = is_con_app e
is_con_app (Tick _ e) = is_con_app e
is_con_app _ = False
\end{code}
Note [Complexity of loop breaking]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The loop-breaking algorithm knocks out one binder at a time, and
performs a new SCC analysis on the remaining binders. That can
behave very badly in tightly-coupled groups of bindings; in the
worst case it can be (N**2)*log N, because it does a full SCC
on N, then N-1, then N-2 and so on.
To avoid this, we switch plans after 2 (or whatever) attempts:
Plan A: pick one binder with the lowest score, make it
a loop breaker, and try again
Plan B: pick *all* binders with the lowest score, make them
all loop breakers, and try again
Since there are only a small finite number of scores, this will
terminate in a constant number of iterations, rather than O(N)
iterations.
You might thing that it's very unlikely, but RULES make it much
more likely. Here's a real example from Trac #1969:
Rec { $dm = \d.\x. op d
{-# RULES forall d. $dm Int d = $s$dm1
forall d. $dm Bool d = $s$dm2 #-}
dInt = MkD .... opInt ...
dInt = MkD .... opBool ...
opInt = $dm dInt
opBool = $dm dBool
$s$dm1 = \x. op dInt
$s$dm2 = \x. op dBool }
The RULES stuff means that we can't choose $dm as a loop breaker
(Note [Choosing loop breakers]), so we must choose at least (say)
opInt *and* opBool, and so on. The number of loop breakders is
linear in the number of instance declarations.
Note [INLINE pragmas]
~~~~~~~~~~~~~~~~~~~~~
Avoid choosing a function with an INLINE pramga as the loop breaker!
If such a function is mutually-recursive with a non-INLINE thing,
then the latter should be the loop-breaker.
----> Historical note, dating from when strictness wrappers
were generated from the strictness signatures:
Usually this is just a question of optimisation. But a particularly
bad case is wrappers generated by the demand analyser: if you make
then into a loop breaker you may get an infinite inlining loop. For
example:
rec {
$wfoo x = ....foo x....
{-loop brk-} foo x = ...$wfoo x...
}
The interface file sees the unfolding for $wfoo, and sees that foo is
strict (and hence it gets an auto-generated wrapper). Result: an
infinite inlining in the importing scope. So be a bit careful if you
change this. A good example is Tree.repTree in
nofib/spectral/minimax. If the repTree wrapper is chosen as the loop
breaker then compiling Game.hs goes into an infinite loop. This
happened when we gave is_con_app a lower score than inline candidates:
Tree.repTree
= __inline_me (/\a. \w w1 w2 ->
case Tree.$wrepTree @ a w w1 w2 of
{ (# ww1, ww2 #) -> Branch @ a ww1 ww2 })
Tree.$wrepTree
= /\a w w1 w2 ->
(# w2_smP, map a (Tree a) (Tree.repTree a w1 w) (w w2) #)
Here we do *not* want to choose 'repTree' as the loop breaker.
-----> End of historical note
Note [DFuns should not be loop breakers]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
It's particularly bad to make a DFun into a loop breaker. See
Note [How instance declarations are translated] in TcInstDcls
We give DFuns a higher score than ordinary CONLIKE things because
if there's a choice we want the DFun to be the non-looop breker. Eg
rec { sc = /\ a \$dC. $fBWrap (T a) ($fCT @ a $dC)
$fCT :: forall a_afE. (Roman.C a_afE) => Roman.C (Roman.T a_afE)
{-# DFUN #-}
$fCT = /\a \$dC. MkD (T a) ((sc @ a $dC) |> blah) ($ctoF @ a $dC)
}
Here 'sc' (the superclass) looks CONLIKE, but we'll never get to it
if we can't unravel the DFun first.
Note [Constructor applications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
It's really really important to inline dictionaries. Real
example (the Enum Ordering instance from GHC.Base):
rec f = \ x -> case d of (p,q,r) -> p x
g = \ x -> case d of (p,q,r) -> q x
d = (v, f, g)
Here, f and g occur just once; but we can't inline them into d.
On the other hand we *could* simplify those case expressions if
we didn't stupidly choose d as the loop breaker.
But we won't because constructor args are marked "Many".
Inlining dictionaries is really essential to unravelling
the loops in static numeric dictionaries, see GHC.Float.
Note [Closure conversion]
~~~~~~~~~~~~~~~~~~~~~~~~~
We treat (\x. C p q) as a high-score candidate in the letrec scoring algorithm.
The immediate motivation came from the result of a closure-conversion transformation
which generated code like this:
data Clo a b = forall c. Clo (c -> a -> b) c
($:) :: Clo a b -> a -> b
Clo f env $: x = f env x
rec { plus = Clo plus1 ()
; plus1 _ n = Clo plus2 n
; plus2 Zero n = n
; plus2 (Succ m) n = Succ (plus $: m $: n) }
If we inline 'plus' and 'plus1', everything unravels nicely. But if
we choose 'plus1' as the loop breaker (which is entirely possible
otherwise), the loop does not unravel nicely.
@occAnalRhs@ deals with the question of bindings where the Id is marked
by an INLINE pragma. For these we record that anything which occurs
in its RHS occurs many times. This pessimistically assumes that ths
inlined binder also occurs many times in its scope, but if it doesn't
we'll catch it next time round. At worst this costs an extra simplifier pass.
ToDo: try using the occurrence info for the inline'd binder.
[March 97] We do the same for atomic RHSs. Reason: see notes with loopBreakSCC.
[June 98, SLPJ] I've undone this change; I don't understand it. See notes with loopBreakSCC.
\begin{code}
occAnalRecRhs :: OccEnv -> CoreExpr
-> (UsageDetails, CoreExpr)
occAnalRecRhs env rhs = occAnal (rhsCtxt env) rhs
occAnalNonRecRhs :: OccEnv
-> Id -> CoreExpr
-> (UsageDetails, CoreExpr)
occAnalNonRecRhs env bndr rhs
= occAnal rhs_env rhs
where
env1 = env { occ_one_shots = argOneShots OneShotLam dmd }
rhs_env | certainly_inline = env1
| otherwise = rhsCtxt env1
certainly_inline
= case idOccInfo bndr of
OneOcc in_lam one_br _ -> not in_lam && one_br && active && not_stable
_ -> False
dmd = idDemandInfo bndr
active = isAlwaysActive (idInlineActivation bndr)
not_stable = not (isStableUnfolding (idUnfolding bndr))
addIdOccs :: UsageDetails -> VarSet -> UsageDetails
addIdOccs usage id_set = foldVarSet add usage id_set
where
add v u | isId v = addOneOcc u v NoOccInfo
| otherwise = u
\end{code}
Note [Cascading inlines]
~~~~~~~~~~~~~~~~~~~~~~~~
By default we use an rhsCtxt for the RHS of a binding. This tells the
occ anal n that it's looking at an RHS, which has an effect in
occAnalApp. In particular, for constructor applications, it makes
the arguments appear to have NoOccInfo, so that we don't inline into
them. Thus x = f y
k = Just x
we do not want to inline x.
But there's a problem. Consider
x1 = a0 : []
x2 = a1 : x1
x3 = a2 : x2
g = f x3
First time round, it looks as if x1 and x2 occur as an arg of a
let-bound constructor ==> give them a many-occurrence.
But then x3 is inlined (unconditionally as it happens) and
next time round, x2 will be, and the next time round x1 will be
Result: multiple simplifier iterations. Sigh.
So, when analysing the RHS of x3 we notice that x3 will itself
definitely inline the next time round, and so we analyse x3's rhs in
an ordinary context, not rhsCtxt. Hence the "certainly_inline" stuff.
Annoyingly, we have to approximate SimplUtils.preInlineUnconditionally.
If we say "yes" when preInlineUnconditionally says "no" the simplifier iterates
indefinitely:
x = f y
k = Just x
inline ==>
k = Just (f y)
float ==>
x1 = f y
k = Just x1
This is worse than the slow cascade, so we only want to say "certainly_inline"
if it really is certain. Look at the note with preInlineUnconditionally
for the various clauses.
Expressions
~~~~~~~~~~~
\begin{code}
occAnal :: OccEnv
-> CoreExpr
-> (UsageDetails,
CoreExpr)
occAnal _ expr@(Type _) = (emptyDetails, expr)
occAnal _ expr@(Lit _) = (emptyDetails, expr)
occAnal env expr@(Var v) = (mkOneOcc env v False, expr)
occAnal _ (Coercion co)
= (addIdOccs emptyDetails (coVarsOfCo co), Coercion co)
\end{code}
Note [Gather occurrences of coercion veriables]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We need to gather info about what coercion variables appear, so that
we can sort them into the right place when doing dependency analysis.
\begin{code}
occAnal env (Tick tickish body)
| Breakpoint _ ids <- tickish
= (mapVarEnv markInsideSCC usage
+++ mkVarEnv (zip ids (repeat NoOccInfo)), Tick tickish body')
| tickishScoped tickish
= (mapVarEnv markInsideSCC usage, Tick tickish body')
| otherwise
= (usage, Tick tickish body')
where
!(usage,body') = occAnal env body
occAnal env (Cast expr co)
= case occAnal env expr of { (usage, expr') ->
let usage1 = markManyIf (isRhsEnv env) usage
usage2 = addIdOccs usage1 (coVarsOfCo co)
in (usage2, Cast expr' co)
}
\end{code}
\begin{code}
occAnal env app@(App _ _)
= occAnalApp env (collectArgs app)
occAnal env (Lam x body) | isTyVar x
= case occAnal env body of { (body_usage, body') ->
(body_usage, Lam x body')
}
occAnal env expr@(Lam _ _)
= case occAnal env_body body of { (body_usage, body') ->
let
(final_usage, tagged_binders) = tagLamBinders body_usage binders'
really_final_usage
| all isOneShotBndr binders' = final_usage
| otherwise = mapVarEnv markInsideLam final_usage
in
(really_final_usage, mkLams tagged_binders body') }
where
(binders, body) = collectBinders expr
(env_body, binders') = oneShotGroup env binders
occAnal env (Case scrut bndr ty alts)
= case occ_anal_scrut scrut alts of { (scrut_usage, scrut') ->
case mapAndUnzip occ_anal_alt alts of { (alts_usage_s, alts') ->
let
alts_usage = foldr combineAltsUsageDetails emptyDetails alts_usage_s
(alts_usage1, tagged_bndr) = tag_case_bndr alts_usage bndr
total_usage = scrut_usage +++ alts_usage1
in
total_usage `seq` (total_usage, Case scrut' tagged_bndr ty alts') }}
where
tag_case_bndr usage bndr
= case lookupVarEnv usage bndr of
Nothing -> (usage, setIdOccInfo bndr IAmDead)
Just _ -> (usage `delVarEnv` bndr, setIdOccInfo bndr NoOccInfo)
alt_env = mkAltEnv env scrut bndr
occ_anal_alt = occAnalAlt alt_env bndr
occ_anal_scrut (Var v) (alt1 : other_alts)
| not (null other_alts) || not (isDefaultAlt alt1)
= (mkOneOcc env v True, Var v)
occ_anal_scrut scrut _alts
= occAnal (vanillaCtxt env) scrut
occAnal env (Let bind body)
= case occAnal env body of { (body_usage, body') ->
case occAnalBind env env emptyVarEnv bind body_usage of { (final_usage, new_binds) ->
(final_usage, mkLets new_binds body') }}
occAnalArgs :: OccEnv -> [CoreExpr] -> [OneShots] -> (UsageDetails, [CoreExpr])
occAnalArgs _ [] _
= (emptyDetails, [])
occAnalArgs env (arg:args) one_shots
| isTypeArg arg
= case occAnalArgs env args one_shots of { (uds, args') ->
(uds, arg:args') }
| otherwise
= case argCtxt env one_shots of { (arg_env, one_shots') ->
case occAnal arg_env arg of { (uds1, arg') ->
case occAnalArgs env args one_shots' of { (uds2, args') ->
(uds1 +++ uds2, arg':args') }}}
\end{code}
Applications are dealt with specially because we want
the "build hack" to work.
Note [Arguments of let-bound constructors]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider
f x = let y = expensive x in
let z = (True,y) in
(case z of {(p,q)->q}, case z of {(p,q)->q})
We feel free to duplicate the WHNF (True,y), but that means
that y may be duplicated thereby.
If we aren't careful we duplicate the (expensive x) call!
Constructors are rather like lambdas in this way.
\begin{code}
occAnalApp :: OccEnv
-> (Expr CoreBndr, [Arg CoreBndr])
-> (UsageDetails, Expr CoreBndr)
occAnalApp env (Var fun, args)
= case args_stuff of { (args_uds, args') ->
let
final_args_uds = markManyIf (isRhsEnv env && is_exp) args_uds
in
(fun_uds +++ final_args_uds, mkApps (Var fun) args') }
where
n_val_args = valArgCount args
fun_uds = mkOneOcc env fun (n_val_args > 0)
is_exp = isExpandableApp fun n_val_args
one_shots = argsOneShots (idStrictness fun) n_val_args
args_stuff = occAnalArgs env args one_shots
occAnalApp env (fun, args)
= case occAnal (addAppCtxt env args) fun of { (fun_uds, fun') ->
case occAnalArgs env args [] of { (args_uds, args') ->
(fun_uds +++ args_uds, mkApps fun' args') }}
markManyIf :: Bool
-> UsageDetails
-> UsageDetails
markManyIf True uds = mapVarEnv markMany uds
markManyIf False uds = uds
\end{code}
Note [Use one-shot information]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The occurrrence analyser propagates one-shot-lambda information in two situation
* Applications: eg build (\cn -> blah)
Propagate one-shot info from the strictness signature of 'build' to
the \cn
* Let-bindings: eg let f = \c. let ... in \n -> blah
in (build f, build f)
Propagate one-shot info from the demanand-info on 'f' to the
lambdas in its RHS (which may not be syntactically at the top)
Some of this is done by the demand analyser, but this way it happens
much earlier, taking advantage of the strictness signature of
imported functions.
Note [Binders in case alternatives]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider
case x of y { (a,b) -> f y }
We treat 'a', 'b' as dead, because they don't physically occur in the
case alternative. (Indeed, a variable is dead iff it doesn't occur in
its scope in the output of OccAnal.) It really helps to know when
binders are unused. See esp the call to isDeadBinder in
Simplify.mkDupableAlt
In this example, though, the Simplifier will bring 'a' and 'b' back to
life, beause it binds 'y' to (a,b) (imagine got inlined and
scrutinised y).
\begin{code}
occAnalAlt :: (OccEnv, Maybe (Id, CoreExpr))
-> CoreBndr
-> CoreAlt
-> (UsageDetails, Alt IdWithOccInfo)
occAnalAlt (env, scrut_bind) case_bndr (con, bndrs, rhs)
= case occAnal env rhs of { (rhs_usage1, rhs1) ->
let
(rhs_usage2, rhs2) =
wrapProxy (occ_binder_swap env) scrut_bind case_bndr rhs_usage1 rhs1
(alt_usg, tagged_bndrs) = tagLamBinders rhs_usage2 bndrs
bndrs' = tagged_bndrs
in
(alt_usg, (con, bndrs', rhs2)) }
wrapProxy :: Bool -> Maybe (Id, CoreExpr) -> Id -> UsageDetails -> CoreExpr -> (UsageDetails, CoreExpr)
wrapProxy enable_binder_swap (Just (scrut_var, rhs)) case_bndr body_usg body
| enable_binder_swap,
scrut_var `usedIn` body_usg
= ( body_usg' +++ unitVarEnv case_bndr NoOccInfo
, Let (NonRec tagged_scrut_var rhs) body )
where
(body_usg', tagged_scrut_var) = tagBinder body_usg scrut_var
wrapProxy _ _ _ body_usg body
= (body_usg, body)
\end{code}
%************************************************************************
%* *
OccEnv
%* *
%************************************************************************
\begin{code}
data OccEnv
= OccEnv { occ_encl :: !OccEncl
, occ_one_shots :: !OneShots
, occ_gbl_scrut :: GlobalScruts
, occ_rule_act :: Activation -> Bool
, occ_binder_swap :: !Bool
}
type GlobalScruts = IdSet
data OccEncl
= OccRhs
| OccVanilla
instance Outputable OccEncl where
ppr OccRhs = ptext (sLit "occRhs")
ppr OccVanilla = ptext (sLit "occVanilla")
type OneShots = [OneShotInfo]
initOccEnv :: (Activation -> Bool) -> OccEnv
initOccEnv active_rule
= OccEnv { occ_encl = OccVanilla
, occ_one_shots = []
, occ_gbl_scrut = emptyVarSet
, occ_rule_act = active_rule
, occ_binder_swap = True }
vanillaCtxt :: OccEnv -> OccEnv
vanillaCtxt env = env { occ_encl = OccVanilla, occ_one_shots = [] }
rhsCtxt :: OccEnv -> OccEnv
rhsCtxt env = env { occ_encl = OccRhs, occ_one_shots = [] }
argCtxt :: OccEnv -> [OneShots] -> (OccEnv, [OneShots])
argCtxt env []
= (env { occ_encl = OccVanilla, occ_one_shots = [] }, [])
argCtxt env (one_shots:one_shots_s)
= (env { occ_encl = OccVanilla, occ_one_shots = one_shots }, one_shots_s)
isRhsEnv :: OccEnv -> Bool
isRhsEnv (OccEnv { occ_encl = OccRhs }) = True
isRhsEnv (OccEnv { occ_encl = OccVanilla }) = False
oneShotGroup :: OccEnv -> [CoreBndr]
-> ( OccEnv
, [CoreBndr] )
oneShotGroup env@(OccEnv { occ_one_shots = ctxt }) bndrs
= go ctxt bndrs []
where
go ctxt [] rev_bndrs
= ( env { occ_one_shots = ctxt, occ_encl = OccVanilla }
, reverse rev_bndrs )
go [] bndrs rev_bndrs
= ( env { occ_one_shots = [], occ_encl = OccVanilla }
, reverse rev_bndrs ++ bndrs )
go ctxt (bndr:bndrs) rev_bndrs
| isId bndr
= case ctxt of
[] -> go [] bndrs (bndr : rev_bndrs)
(one_shot : ctxt) -> go ctxt bndrs (bndr': rev_bndrs)
where
bndr' = updOneShotInfo bndr one_shot
| otherwise
= go ctxt bndrs (bndr:rev_bndrs)
addAppCtxt :: OccEnv -> [Arg CoreBndr] -> OccEnv
addAppCtxt env@(OccEnv { occ_one_shots = ctxt }) args
= env { occ_one_shots = replicate (valArgCount args) OneShotLam ++ ctxt }
\end{code}
\begin{code}
transClosureFV :: UniqFM VarSet -> UniqFM VarSet
transClosureFV env
| no_change = env
| otherwise = transClosureFV (listToUFM new_fv_list)
where
(no_change, new_fv_list) = mapAccumL bump True (ufmToList env)
bump no_change (b,fvs)
| no_change_here = (no_change, (b,fvs))
| otherwise = (False, (b,new_fvs))
where
(new_fvs, no_change_here) = extendFvs env fvs
extendFvs_ :: UniqFM VarSet -> VarSet -> VarSet
extendFvs_ env s = fst (extendFvs env s)
extendFvs :: UniqFM VarSet -> VarSet -> (VarSet, Bool)
extendFvs env s
| isNullUFM env
= (s, True)
| otherwise
= (s `unionVarSet` extras, extras `subVarSet` s)
where
extras :: VarSet
extras = foldUFM unionVarSet emptyVarSet $
intersectUFM_C (\x _ -> x) env s
\end{code}
%************************************************************************
%* *
Binder swap
%* *
%************************************************************************
Note [Binder swap]
~~~~~~~~~~~~~~~~~~
We do these two transformations right here:
(1) case x of b { pi -> ri }
==>
case x of b { pi -> let x=b in ri }
(2) case (x |> co) of b { pi -> ri }
==>
case (x |> co) of b { pi -> let x = b |> sym co in ri }
Why (2)? See Note [Case of cast]
In both cases, in a particular alternative (pi -> ri), we only
add the binding if
(a) x occurs free in (pi -> ri)
(ie it occurs in ri, but is not bound in pi)
(b) the pi does not bind b (or the free vars of co)
We need (a) and (b) for the inserted binding to be correct.
For the alternatives where we inject the binding, we can transfer
all x's OccInfo to b. And that is the point.
Notice that
* The deliberate shadowing of 'x'.
* That (a) rapidly becomes false, so no bindings are injected.
The reason for doing these transformations here is because it allows
us to adjust the OccInfo for 'x' and 'b' as we go.
* Suppose the only occurrences of 'x' are the scrutinee and in the
ri; then this transformation makes it occur just once, and hence
get inlined right away.
* If we do this in the Simplifier, we don't know whether 'x' is used
in ri, so we are forced to pessimistically zap b's OccInfo even
though it is typically dead (ie neither it nor x appear in the
ri). There's nothing actually wrong with zapping it, except that
it's kind of nice to know which variables are dead. My nose
tells me to keep this information as robustly as possible.
The Maybe (Id,CoreExpr) passed to occAnalAlt is the extra let-binding
{x=b}; it's Nothing if the binder-swap doesn't happen.
There is a danger though. Consider
let v = x +# y
in case (f v) of w -> ...v...v...
And suppose that (f v) expands to just v. Then we'd like to
use 'w' instead of 'v' in the alternative. But it may be too
late; we may have substituted the (cheap) x+#y for v in the
same simplifier pass that reduced (f v) to v.
I think this is just too bad. CSE will recover some of it.
Note [Case of cast]
~~~~~~~~~~~~~~~~~~~
Consider case (x `cast` co) of b { I# ->
... (case (x `cast` co) of {...}) ...
We'd like to eliminate the inner case. That is the motivation for
equation (2) in Note [Binder swap]. When we get to the inner case, we
inline x, cancel the casts, and away we go.
Note [Binder swap on GlobalId scrutinees]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When the scrutinee is a GlobalId we must take care in two ways
i) In order to *know* whether 'x' occurs free in the RHS, we need its
occurrence info. BUT, we don't gather occurrence info for
GlobalIds. That's the reason for the (small) occ_gbl_scrut env in
OccEnv is for: it says "gather occurrence info for these".
ii) We must call localiseId on 'x' first, in case it's a GlobalId, or
has an External Name. See, for example, SimplEnv Note [Global Ids in
the substitution].
Note [Zap case binders in proxy bindings]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
From the original
case x of cb(dead) { p -> ...x... }
we will get
case x of cb(live) { p -> let x = cb in ...x... }
Core Lint never expects to find an *occurrence* of an Id marked
as Dead, so we must zap the OccInfo on cb before making the
binding x = cb. See Trac #5028.
Historical note [no-case-of-case]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We *used* to suppress the binder-swap in case expressions when
-fno-case-of-case is on. Old remarks:
"This happens in the first simplifier pass,
and enhances full laziness. Here's the bad case:
f = \ y -> ...(case x of I# v -> ...(case x of ...) ... )
If we eliminate the inner case, we trap it inside the I# v -> arm,
which might prevent some full laziness happening. I've seen this
in action in spectral/cichelli/Prog.hs:
[(m,n) | m <- [1..max], n <- [1..max]]
Hence the check for NoCaseOfCase."
However, now the full-laziness pass itself reverses the binder-swap, so this
check is no longer necessary.
Historical note [Suppressing the case binder-swap]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
This old note describes a problem that is also fixed by doing the
binder-swap in OccAnal:
There is another situation when it might make sense to suppress the
case-expression binde-swap. If we have
case x of w1 { DEFAULT -> case x of w2 { A -> e1; B -> e2 }
...other cases .... }
We'll perform the binder-swap for the outer case, giving
case x of w1 { DEFAULT -> case w1 of w2 { A -> e1; B -> e2 }
...other cases .... }
But there is no point in doing it for the inner case, because w1 can't
be inlined anyway. Furthermore, doing the case-swapping involves
zapping w2's occurrence info (see paragraphs that follow), and that
forces us to bind w2 when doing case merging. So we get
case x of w1 { A -> let w2 = w1 in e1
B -> let w2 = w1 in e2
...other cases .... }
This is plain silly in the common case where w2 is dead.
Even so, I can't see a good way to implement this idea. I tried
not doing the binder-swap if the scrutinee was already evaluated
but that failed big-time:
data T = MkT !Int
case v of w { MkT x ->
case x of x1 { I# y1 ->
case x of x2 { I# y2 -> ...
Notice that because MkT is strict, x is marked "evaluated". But to
eliminate the last case, we must either make sure that x (as well as
x1) has unfolding MkT y1. THe straightforward thing to do is to do
the binder-swap. So this whole note is a no-op.
It's fixed by doing the binder-swap in OccAnal because we can do the
binder-swap unconditionally and still get occurrence analysis
information right.
\begin{code}
mkAltEnv :: OccEnv -> CoreExpr -> Id -> (OccEnv, Maybe (Id, CoreExpr))
mkAltEnv env@(OccEnv { occ_gbl_scrut = pe }) scrut case_bndr
= case scrut of
Var v -> add_scrut v case_bndr'
Cast (Var v) co -> add_scrut v (Cast case_bndr' (mkSymCo co))
_ -> (env { occ_encl = OccVanilla }, Nothing)
where
add_scrut v rhs = ( env { occ_encl = OccVanilla, occ_gbl_scrut = pe `extendVarSet` v }
, Just (localise v, rhs) )
case_bndr' = Var (zapIdOccInfo case_bndr)
localise scrut_var = mkLocalId (localiseName (idName scrut_var)) (idType scrut_var)
\end{code}
%************************************************************************
%* *
\subsection[OccurAnal-types]{OccEnv}
%* *
%************************************************************************
\begin{code}
type UsageDetails = IdEnv OccInfo
(+++), combineAltsUsageDetails
:: UsageDetails -> UsageDetails -> UsageDetails
(+++) usage1 usage2
= plusVarEnv_C addOccInfo usage1 usage2
combineAltsUsageDetails usage1 usage2
= plusVarEnv_C orOccInfo usage1 usage2
addOneOcc :: UsageDetails -> Id -> OccInfo -> UsageDetails
addOneOcc usage id info
= plusVarEnv_C addOccInfo usage (unitVarEnv id info)
emptyDetails :: UsageDetails
emptyDetails = (emptyVarEnv :: UsageDetails)
usedIn :: Id -> UsageDetails -> Bool
v `usedIn` details = isExportedId v || v `elemVarEnv` details
type IdWithOccInfo = Id
tagLamBinders :: UsageDetails
-> [Id]
-> (UsageDetails,
[IdWithOccInfo])
tagLamBinders usage binders = usage' `seq` (usage', bndrs')
where
(usage', bndrs') = mapAccumR tag_lam usage binders
tag_lam usage bndr = (usage2, setBinderOcc usage bndr)
where
usage1 = usage `delVarEnv` bndr
usage2 | isId bndr = addIdOccs usage1 (idUnfoldingVars bndr)
| otherwise = usage1
tagBinder :: UsageDetails
-> Id
-> (UsageDetails,
IdWithOccInfo)
tagBinder usage binder
= let
usage' = usage `delVarEnv` binder
binder' = setBinderOcc usage binder
in
usage' `seq` (usage', binder')
setBinderOcc :: UsageDetails -> CoreBndr -> CoreBndr
setBinderOcc usage bndr
| isTyVar bndr = bndr
| isExportedId bndr = case idOccInfo bndr of
NoOccInfo -> bndr
_ -> setIdOccInfo bndr NoOccInfo
| otherwise = setIdOccInfo bndr occ_info
where
occ_info = lookupVarEnv usage bndr `orElse` IAmDead
\end{code}
%************************************************************************
%* *
\subsection{Operations over OccInfo}
%* *
%************************************************************************
\begin{code}
mkOneOcc :: OccEnv -> Id -> InterestingCxt -> UsageDetails
mkOneOcc env id int_cxt
| isLocalId id
= unitVarEnv id (OneOcc False True int_cxt)
| id `elemVarEnv` occ_gbl_scrut env
= unitVarEnv id NoOccInfo
| otherwise
= emptyDetails
markMany, markInsideLam, markInsideSCC :: OccInfo -> OccInfo
markMany _ = NoOccInfo
markInsideSCC occ = markInsideLam occ
markInsideLam (OneOcc _ one_br int_cxt) = OneOcc True one_br int_cxt
markInsideLam occ = occ
addOccInfo, orOccInfo :: OccInfo -> OccInfo -> OccInfo
addOccInfo a1 a2 = ASSERT( not (isDeadOcc a1 || isDeadOcc a2) )
NoOccInfo
orOccInfo (OneOcc in_lam1 _ int_cxt1)
(OneOcc in_lam2 _ int_cxt2)
= OneOcc (in_lam1 || in_lam2)
False
(int_cxt1 && int_cxt2)
orOccInfo a1 a2 = ASSERT( not (isDeadOcc a1 || isDeadOcc a2) )
NoOccInfo
\end{code}