% % (c) The AQUA Project, Glasgow University, 1993-1998 % \section[SimplUtils]{The simplifier utilities} \begin{code} module SimplUtils ( -- Rebuilding mkLam, mkCase, prepareAlts, bindCaseBndr, -- Inlining, preInlineUnconditionally, postInlineUnconditionally, activeInline, activeRule, inlineMode, -- The continuation type SimplCont(..), DupFlag(..), LetRhsFlag(..), contIsDupable, contResultType, contIsTrivial, contArgs, dropArgs, countValArgs, countArgs, mkBoringStop, mkLazyArgStop, mkRhsStop, contIsRhsOrArg, interestingCallContext, interestingArgContext, interestingArg, mkArgInfo ) where #include "HsVersions.h" import SimplEnv import DynFlags import StaticFlags import CoreSyn import PprCore import CoreFVs import CoreUtils import Literal import CoreUnfold import MkId import Id import NewDemand import SimplMonad import Type import TyCon import DataCon import TcGadt ( dataConCanMatch ) import VarSet import BasicTypes import Util import Outputable import List( nub ) \end{code} %************************************************************************ %* * The SimplCont type %* * %************************************************************************ A SimplCont allows the simplifier to traverse the expression in a zipper-like fashion. The SimplCont represents the rest of the expression, "above" the point of interest. You can also think of a SimplCont as an "evaluation context", using that term in the way it is used for operational semantics. This is the way I usually think of it, For example you'll often see a syntax for evaluation context looking like C ::= [] | C e | case C of alts | C `cast` co That's the kind of thing we are doing here, and I use that syntax in the comments. Key points: * A SimplCont describes a *strict* context (just like evaluation contexts do). E.g. Just [] is not a SimplCont * A SimplCont describes a context that *does not* bind any variables. E.g. \x. [] is not a SimplCont \begin{code} data SimplCont = Stop -- An empty context, or hole, [] OutType -- Type of the result LetRhsFlag Bool -- True <=> There is something interesting about -- the context, and hence the inliner -- should be a bit keener (see interestingCallContext) -- Two cases: -- (a) This is the RHS of a thunk whose type suggests -- that update-in-place would be possible -- (b) This is an argument of a function that has RULES -- Inlining the call might allow the rule to fire | CoerceIt -- C `cast` co OutCoercion -- The coercion simplified SimplCont | ApplyTo -- C arg DupFlag InExpr SimplEnv -- The argument and its static env SimplCont | Select -- case C of alts DupFlag InId [InAlt] SimplEnv -- The case binder, alts, and subst-env SimplCont -- The two strict forms have no DupFlag, because we never duplicate them | StrictBind -- (\x* \xs. e) C InId [InBndr] -- let x* = [] in e InExpr SimplEnv -- is a special case SimplCont | StrictArg -- e C OutExpr OutType -- e and its type (Bool,[Bool]) -- Whether the function at the head of e has rules, SimplCont -- plus strictness flags for further args data LetRhsFlag = AnArg -- It's just an argument not a let RHS | AnRhs -- It's the RHS of a let (so please float lets out of big lambdas) instance Outputable LetRhsFlag where ppr AnArg = ptext SLIT("arg") ppr AnRhs = ptext SLIT("rhs") instance Outputable SimplCont where ppr (Stop ty is_rhs _) = ptext SLIT("Stop") <> brackets (ppr is_rhs) <+> ppr ty ppr (ApplyTo dup arg se cont) = ((ptext SLIT("ApplyTo") <+> ppr dup <+> pprParendExpr arg) $$ nest 2 (pprSimplEnv se)) $$ ppr cont ppr (StrictBind b _ _ _ cont) = (ptext SLIT("StrictBind") <+> ppr b) $$ ppr cont ppr (StrictArg f _ _ cont) = (ptext SLIT("StrictArg") <+> ppr f) $$ ppr cont ppr (Select dup bndr alts se cont) = (ptext SLIT("Select") <+> ppr dup <+> ppr bndr) $$ (nest 4 (ppr alts $$ pprSimplEnv se)) $$ ppr cont ppr (CoerceIt co cont) = (ptext SLIT("CoerceIt") <+> ppr co) $$ ppr cont data DupFlag = OkToDup | NoDup instance Outputable DupFlag where ppr OkToDup = ptext SLIT("ok") ppr NoDup = ptext SLIT("nodup") ------------------- mkBoringStop :: OutType -> SimplCont mkBoringStop ty = Stop ty AnArg False mkLazyArgStop :: OutType -> Bool -> SimplCont mkLazyArgStop ty has_rules = Stop ty AnArg (canUpdateInPlace ty || has_rules) mkRhsStop :: OutType -> SimplCont mkRhsStop ty = Stop ty AnRhs (canUpdateInPlace ty) contIsRhsOrArg (Stop _ _ _) = True contIsRhsOrArg (StrictBind {}) = True contIsRhsOrArg (StrictArg {}) = True contIsRhsOrArg other = False ------------------- contIsDupable :: SimplCont -> Bool contIsDupable (Stop _ _ _) = True contIsDupable (ApplyTo OkToDup _ _ _) = True contIsDupable (Select OkToDup _ _ _ _) = True contIsDupable (CoerceIt _ cont) = contIsDupable cont contIsDupable other = False ------------------- contIsTrivial :: SimplCont -> Bool contIsTrivial (Stop _ _ _) = True contIsTrivial (ApplyTo _ (Type _) _ cont) = contIsTrivial cont contIsTrivial (CoerceIt _ cont) = contIsTrivial cont contIsTrivial other = False ------------------- contResultType :: SimplCont -> OutType contResultType (Stop to_ty _ _) = to_ty contResultType (StrictArg _ _ _ cont) = contResultType cont contResultType (StrictBind _ _ _ _ cont) = contResultType cont contResultType (ApplyTo _ _ _ cont) = contResultType cont contResultType (CoerceIt _ cont) = contResultType cont contResultType (Select _ _ _ _ cont) = contResultType cont ------------------- countValArgs :: SimplCont -> Int countValArgs (ApplyTo _ (Type ty) se cont) = countValArgs cont countValArgs (ApplyTo _ val_arg se cont) = 1 + countValArgs cont countValArgs other = 0 countArgs :: SimplCont -> Int countArgs (ApplyTo _ arg se cont) = 1 + countArgs cont countArgs other = 0 contArgs :: SimplCont -> ([OutExpr], SimplCont) -- Uses substitution to turn each arg into an OutExpr contArgs cont = go [] cont where go args (ApplyTo _ arg se cont) = go (substExpr se arg : args) cont go args cont = (reverse args, cont) dropArgs :: Int -> SimplCont -> SimplCont dropArgs 0 cont = cont dropArgs n (ApplyTo _ _ _ cont) = dropArgs (n-1) cont dropArgs n other = pprPanic "dropArgs" (ppr n <+> ppr other) \end{code} \begin{code} interestingArg :: OutExpr -> Bool -- An argument is interesting if it has *some* structure -- We are here trying to avoid unfolding a function that -- is applied only to variables that have no unfolding -- (i.e. they are probably lambda bound): f x y z -- There is little point in inlining f here. interestingArg (Var v) = hasSomeUnfolding (idUnfolding v) -- Was: isValueUnfolding (idUnfolding v') -- But that seems over-pessimistic || isDataConWorkId v -- This accounts for an argument like -- () or [], which is definitely interesting interestingArg (Type _) = False interestingArg (App fn (Type _)) = interestingArg fn interestingArg (Note _ a) = interestingArg a -- Idea (from Sam B); I'm not sure if it's a good idea, so commented out for now -- interestingArg expr | isUnLiftedType (exprType expr) -- -- Unlifted args are only ever interesting if we know what they are -- = case expr of -- Lit lit -> True -- _ -> False interestingArg other = True -- Consider let x = 3 in f x -- The substitution will contain (x -> ContEx 3), and we want to -- to say that x is an interesting argument. -- But consider also (\x. f x y) y -- The substitution will contain (x -> ContEx y), and we want to say -- that x is not interesting (assuming y has no unfolding) \end{code} Comment about interestingCallContext ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We want to avoid inlining an expression where there can't possibly be any gain, such as in an argument position. Hence, if the continuation is interesting (eg. a case scrutinee, application etc.) then we inline, otherwise we don't. Previously some_benefit used to return True only if the variable was applied to some value arguments. This didn't work: let x = _coerce_ (T Int) Int (I# 3) in case _coerce_ Int (T Int) x of I# y -> .... we want to inline x, but can't see that it's a constructor in a case scrutinee position, and some_benefit is False. Another example: dMonadST = _/\_ t -> :Monad (g1 _@_ t, g2 _@_ t, g3 _@_ t) .... case dMonadST _@_ x0 of (a,b,c) -> .... we'd really like to inline dMonadST here, but we *don't* want to inline if the case expression is just case x of y { DEFAULT -> ... } since we can just eliminate this case instead (x is in WHNF). Similar applies when x is bound to a lambda expression. Hence contIsInteresting looks for case expressions with just a single default case. \begin{code} interestingCallContext :: Bool -- False <=> no args at all -> Bool -- False <=> no value args -> SimplCont -> Bool -- The "lone-variable" case is important. I spent ages -- messing about with unsatisfactory varaints, but this is nice. -- The idea is that if a variable appear all alone -- as an arg of lazy fn, or rhs Stop -- as scrutinee of a case Select -- as arg of a strict fn ArgOf -- then we should not inline it (unless there is some other reason, -- e.g. is is the sole occurrence). We achieve this by making -- interestingCallContext return False for a lone variable. -- -- Why? At least in the case-scrutinee situation, turning -- let x = (a,b) in case x of y -> ... -- into -- let x = (a,b) in case (a,b) of y -> ... -- and thence to -- let x = (a,b) in let y = (a,b) in ... -- is bad if the binding for x will remain. -- -- Another example: I discovered that strings -- were getting inlined straight back into applications of 'error' -- because the latter is strict. -- s = "foo" -- f = \x -> ...(error s)... -- Fundamentally such contexts should not ecourage inlining because -- the context can ``see'' the unfolding of the variable (e.g. case or a RULE) -- so there's no gain. -- -- However, even a type application or coercion isn't a lone variable. -- Consider -- case $fMonadST @ RealWorld of { :DMonad a b c -> c } -- We had better inline that sucker! The case won't see through it. -- -- For now, I'm treating treating a variable applied to types -- in a *lazy* context "lone". The motivating example was -- f = /\a. \x. BIG -- g = /\a. \y. h (f a) -- There's no advantage in inlining f here, and perhaps -- a significant disadvantage. Hence some_val_args in the Stop case interestingCallContext some_args some_val_args cont = interesting cont where interesting (Select {}) = some_args interesting (ApplyTo {}) = True -- Can happen if we have (coerce t (f x)) y -- Perhaps True is a bit over-keen, but I've -- seen (coerce f) x, where f has an INLINE prag, -- So we have to give some motivaiton for inlining it interesting (StrictArg {}) = some_val_args interesting (StrictBind {}) = some_val_args -- ?? interesting (Stop ty _ interesting) = some_val_args && interesting interesting (CoerceIt _ cont) = interesting cont -- If this call is the arg of a strict function, the context -- is a bit interesting. If we inline here, we may get useful -- evaluation information to avoid repeated evals: e.g. -- x + (y * z) -- Here the contIsInteresting makes the '*' keener to inline, -- which in turn exposes a constructor which makes the '+' inline. -- Assuming that +,* aren't small enough to inline regardless. -- -- It's also very important to inline in a strict context for things -- like -- foldr k z (f x) -- Here, the context of (f x) is strict, and if f's unfolding is -- a build it's *great* to inline it here. So we must ensure that -- the context for (f x) is not totally uninteresting. ------------------- mkArgInfo :: Id -> Int -- Number of value args -> SimplCont -- Context of the cal -> (Bool, [Bool]) -- Arg info -- The arg info consists of -- * A Bool indicating if the function has rules (recursively) -- * A [Bool] indicating strictness for each arg -- The [Bool] is usually infinite, but if it is finite it -- guarantees that the function diverges after being given -- that number of args mkArgInfo fun n_val_args call_cont = (interestingArgContext fun call_cont, fun_stricts) where vanilla_stricts, fun_stricts :: [Bool] vanilla_stricts = repeat False fun_stricts = case splitStrictSig (idNewStrictness fun) of (demands, result_info) | not (demands `lengthExceeds` n_val_args) -> -- Enough args, use the strictness given. -- For bottoming functions we used to pretend that the arg -- is lazy, so that we don't treat the arg as an -- interesting context. This avoids substituting -- top-level bindings for (say) strings into -- calls to error. But now we are more careful about -- inlining lone variables, so its ok (see SimplUtils.analyseCont) if isBotRes result_info then map isStrictDmd demands -- Finite => result is bottom else map isStrictDmd demands ++ vanilla_stricts other -> vanilla_stricts -- Not enough args, or no strictness interestingArgContext :: Id -> SimplCont -> Bool -- If the argument has form (f x y), where x,y are boring, -- and f is marked INLINE, then we don't want to inline f. -- But if the context of the argument is -- g (f x y) -- where g has rules, then we *do* want to inline f, in case it -- exposes a rule that might fire. Similarly, if the context is -- h (g (f x x)) -- where h has rules, then we do want to inline f. -- The interesting_arg_ctxt flag makes this happen; if it's -- set, the inliner gets just enough keener to inline f -- regardless of how boring f's arguments are, if it's marked INLINE -- -- The alternative would be to *always* inline an INLINE function, -- regardless of how boring its context is; but that seems overkill -- For example, it'd mean that wrapper functions were always inlined interestingArgContext fn cont = idHasRules fn || go cont where go (Select {}) = False go (ApplyTo {}) = False go (StrictArg {}) = True go (StrictBind {}) = False -- ?? go (CoerceIt _ c) = go c go (Stop _ _ interesting) = interesting ------------------- canUpdateInPlace :: Type -> Bool -- Consider let x = in ... -- If returns an explicit constructor, we might be able -- to do update in place. So we treat even a thunk RHS context -- as interesting if update in place is possible. We approximate -- this by seeing if the type has a single constructor with a -- small arity. But arity zero isn't good -- we share the single copy -- for that case, so no point in sharing. canUpdateInPlace ty | not opt_UF_UpdateInPlace = False | otherwise = case splitTyConApp_maybe ty of Nothing -> False Just (tycon, _) -> case tyConDataCons_maybe tycon of Just [dc] -> arity == 1 || arity == 2 where arity = dataConRepArity dc other -> False \end{code} %************************************************************************ %* * \subsection{Decisions about inlining} %* * %************************************************************************ Inlining is controlled partly by the SimplifierMode switch. This has two settings: SimplGently (a) Simplifying before specialiser/full laziness (b) Simplifiying inside INLINE pragma (c) Simplifying the LHS of a rule (d) Simplifying a GHCi expression or Template Haskell splice SimplPhase n Used at all other times The key thing about SimplGently is that it does no call-site inlining. Before full laziness we must be careful not to inline wrappers, because doing so inhibits floating e.g. ...(case f x of ...)... ==> ...(case (case x of I# x# -> fw x#) of ...)... ==> ...(case x of I# x# -> case fw x# of ...)... and now the redex (f x) isn't floatable any more. The no-inlining thing is also important for Template Haskell. You might be compiling in one-shot mode with -O2; but when TH compiles a splice before running it, we don't want to use -O2. Indeed, we don't want to inline anything, because the byte-code interpreter might get confused about unboxed tuples and suchlike. INLINE pragmas ~~~~~~~~~~~~~~ SimplGently is also used as the mode to simplify inside an InlineMe note. \begin{code} inlineMode :: SimplifierMode inlineMode = SimplGently \end{code} It really is important to switch off inlinings inside such expressions. Consider the following example let f = \pq -> BIG in let g = \y -> f y y {-# INLINE g #-} in ...g...g...g...g...g... Now, if that's the ONLY occurrence of f, it will be inlined inside g, and thence copied multiple times when g is inlined. This function may be inlinined in other modules, so we don't want to remove (by inlining) calls to functions that have specialisations, or that may have transformation rules in an importing scope. E.g. {-# INLINE f #-} f x = ...g... and suppose that g is strict *and* has specialisations. If we inline g's wrapper, we deny f the chance of getting the specialised version of g when f is inlined at some call site (perhaps in some other module). It's also important not to inline a worker back into a wrapper. A wrapper looks like wraper = inline_me (\x -> ...worker... ) Normally, the inline_me prevents the worker getting inlined into the wrapper (initially, the worker's only call site!). But, if the wrapper is sure to be called, the strictness analyser will mark it 'demanded', so when the RHS is simplified, it'll get an ArgOf continuation. That's why the keep_inline predicate returns True for ArgOf continuations. It shouldn't do any harm not to dissolve the inline-me note under these circumstances. Note that the result is that we do very little simplification inside an InlineMe. all xs = foldr (&&) True xs any p = all . map p {-# INLINE any #-} Problem: any won't get deforested, and so if it's exported and the importer doesn't use the inlining, (eg passes it as an arg) then we won't get deforestation at all. We havn't solved this problem yet! preInlineUnconditionally ~~~~~~~~~~~~~~~~~~~~~~~~ @preInlineUnconditionally@ examines a bndr to see if it is used just once in a completely safe way, so that it is safe to discard the binding inline its RHS at the (unique) usage site, REGARDLESS of how big the RHS might be. If this is the case we don't simplify the RHS first, but just inline it un-simplified. This is much better than first simplifying a perhaps-huge RHS and then inlining and re-simplifying it. Indeed, it can be at least quadratically better. Consider x1 = e1 x2 = e2[x1] x3 = e3[x2] ...etc... xN = eN[xN-1] We may end up simplifying e1 N times, e2 N-1 times, e3 N-3 times etc. This can happen with cascades of functions too: f1 = \x1.e1 f2 = \xs.e2[f1] f3 = \xs.e3[f3] ...etc... THE MAIN INVARIANT is this: ---- preInlineUnconditionally invariant ----- IF preInlineUnconditionally chooses to inline x = THEN doing the inlining should not change the occurrence info for the free vars of ---------------------------------------------- For example, it's tempting to look at trivial binding like x = y and inline it unconditionally. But suppose x is used many times, but this is the unique occurrence of y. Then inlining x would change y's occurrence info, which breaks the invariant. It matters: y might have a BIG rhs, which will now be dup'd at every occurrenc of x. Evne RHSs labelled InlineMe aren't caught here, because there might be no benefit from inlining at the call site. [Sept 01] Don't unconditionally inline a top-level thing, because that can simply make a static thing into something built dynamically. E.g. x = (a,b) main = \s -> h x [Remember that we treat \s as a one-shot lambda.] No point in inlining x unless there is something interesting about the call site. But watch out: if you aren't careful, some useful foldr/build fusion can be lost (most notably in spectral/hartel/parstof) because the foldr didn't see the build. Doing the dynamic allocation isn't a big deal, in fact, but losing the fusion can be. But the right thing here seems to be to do a callSiteInline based on the fact that there is something interesting about the call site (it's strict). Hmm. That seems a bit fragile. Conclusion: inline top level things gaily until Phase 0 (the last phase), at which point don't. \begin{code} preInlineUnconditionally :: SimplEnv -> TopLevelFlag -> InId -> InExpr -> Bool preInlineUnconditionally env top_lvl bndr rhs | not active = False | opt_SimplNoPreInlining = False | otherwise = case idOccInfo bndr of IAmDead -> True -- Happens in ((\x.1) v) OneOcc in_lam True int_cxt -> try_once in_lam int_cxt other -> False where phase = getMode env active = case phase of SimplGently -> isAlwaysActive prag SimplPhase n -> isActive n prag prag = idInlinePragma bndr try_once in_lam int_cxt -- There's one textual occurrence | not in_lam = isNotTopLevel top_lvl || early_phase | otherwise = int_cxt && canInlineInLam rhs -- Be very careful before inlining inside a lambda, becuase (a) we must not -- invalidate occurrence information, and (b) we want to avoid pushing a -- single allocation (here) into multiple allocations (inside lambda). -- Inlining a *function* with a single *saturated* call would be ok, mind you. -- || (if is_cheap && not (canInlineInLam rhs) then pprTrace "preinline" (ppr bndr <+> ppr rhs) ok else ok) -- where -- is_cheap = exprIsCheap rhs -- ok = is_cheap && int_cxt -- int_cxt The context isn't totally boring -- E.g. let f = \ab.BIG in \y. map f xs -- Don't want to substitute for f, because then we allocate -- its closure every time the \y is called -- But: let f = \ab.BIG in \y. map (f y) xs -- Now we do want to substitute for f, even though it's not -- saturated, because we're going to allocate a closure for -- (f y) every time round the loop anyhow. -- canInlineInLam => free vars of rhs are (Once in_lam) or Many, -- so substituting rhs inside a lambda doesn't change the occ info. -- Sadly, not quite the same as exprIsHNF. canInlineInLam (Lit l) = True canInlineInLam (Lam b e) = isRuntimeVar b || canInlineInLam e canInlineInLam (Note _ e) = canInlineInLam e canInlineInLam _ = False early_phase = case phase of SimplPhase 0 -> False other -> True -- If we don't have this early_phase test, consider -- x = length [1,2,3] -- The full laziness pass carefully floats all the cons cells to -- top level, and preInlineUnconditionally floats them all back in. -- Result is (a) static allocation replaced by dynamic allocation -- (b) many simplifier iterations because this tickles -- a related problem; only one inlining per pass -- -- On the other hand, I have seen cases where top-level fusion is -- lost if we don't inline top level thing (e.g. string constants) -- Hence the test for phase zero (which is the phase for all the final -- simplifications). Until phase zero we take no special notice of -- top level things, but then we become more leery about inlining -- them. \end{code} postInlineUnconditionally ~~~~~~~~~~~~~~~~~~~~~~~~~ @postInlineUnconditionally@ decides whether to unconditionally inline a thing based on the form of its RHS; in particular if it has a trivial RHS. If so, we can inline and discard the binding altogether. NB: a loop breaker has must_keep_binding = True and non-loop-breakers only have *forward* references Hence, it's safe to discard the binding NOTE: This isn't our last opportunity to inline. We're at the binding site right now, and we'll get another opportunity when we get to the ocurrence(s) Note that we do this unconditional inlining only for trival RHSs. Don't inline even WHNFs inside lambdas; doing so may simply increase allocation when the function is called. This isn't the last chance; see NOTE above. NB: Even inline pragmas (e.g. IMustBeINLINEd) are ignored here Why? Because we don't even want to inline them into the RHS of constructor arguments. See NOTE above NB: At one time even NOINLINE was ignored here: if the rhs is trivial it's best to inline it anyway. We often get a=E; b=a from desugaring, with both a and b marked NOINLINE. But that seems incompatible with our new view that inlining is like a RULE, so I'm sticking to the 'active' story for now. \begin{code} postInlineUnconditionally :: SimplEnv -> TopLevelFlag -> InId -- The binder (an OutId would be fine too) -> OccInfo -- From the InId -> OutExpr -> Unfolding -> Bool postInlineUnconditionally env top_lvl bndr occ_info rhs unfolding | not active = False | isLoopBreaker occ_info = False -- If it's a loop-breaker of any kind, dont' inline -- because it might be referred to "earlier" | isExportedId bndr = False | exprIsTrivial rhs = True | otherwise = case occ_info of -- The point of examining occ_info here is that for *non-values* -- that occur outside a lambda, the call-site inliner won't have -- a chance (becuase it doesn't know that the thing -- only occurs once). The pre-inliner won't have gotten -- it either, if the thing occurs in more than one branch -- So the main target is things like -- let x = f y in -- case v of -- True -> case x of ... -- False -> case x of ... -- I'm not sure how important this is in practice OneOcc in_lam one_br int_cxt -- OneOcc => no code-duplication issue -> smallEnoughToInline unfolding -- Small enough to dup -- ToDo: consider discount on smallEnoughToInline if int_cxt is true -- -- NB: Do NOT inline arbitrarily big things, even if one_br is True -- Reason: doing so risks exponential behaviour. We simplify a big -- expression, inline it, and simplify it again. But if the -- very same thing happens in the big expression, we get -- exponential cost! -- PRINCIPLE: when we've already simplified an expression once, -- make sure that we only inline it if it's reasonably small. && ((isNotTopLevel top_lvl && not in_lam) || -- But outside a lambda, we want to be reasonably aggressive -- about inlining into multiple branches of case -- e.g. let x = -- in case y of { C1 -> ..x..; C2 -> ..x..; C3 -> ... } -- Inlining can be a big win if C3 is the hot-spot, even if -- the uses in C1, C2 are not 'interesting' -- An example that gets worse if you add int_cxt here is 'clausify' (isCheapUnfolding unfolding && int_cxt)) -- isCheap => acceptable work duplication; in_lam may be true -- int_cxt to prevent us inlining inside a lambda without some -- good reason. See the notes on int_cxt in preInlineUnconditionally IAmDead -> True -- This happens; for example, the case_bndr during case of -- known constructor: case (a,b) of x { (p,q) -> ... } -- Here x isn't mentioned in the RHS, so we don't want to -- create the (dead) let-binding let x = (a,b) in ... other -> False -- Here's an example that we don't handle well: -- let f = if b then Left (\x.BIG) else Right (\y.BIG) -- in \y. ....case f of {...} .... -- Here f is used just once, and duplicating the case work is fine (exprIsCheap). -- But -- * We can't preInlineUnconditionally because that woud invalidate -- the occ info for b. -- * We can't postInlineUnconditionally because the RHS is big, and -- that risks exponential behaviour -- * We can't call-site inline, because the rhs is big -- Alas! where active = case getMode env of SimplGently -> isAlwaysActive prag SimplPhase n -> isActive n prag prag = idInlinePragma bndr activeInline :: SimplEnv -> OutId -> Bool activeInline env id = case getMode env of SimplGently -> False -- No inlining at all when doing gentle stuff, -- except for local things that occur once -- The reason is that too little clean-up happens if you -- don't inline use-once things. Also a bit of inlining is *good* for -- full laziness; it can expose constant sub-expressions. -- Example in spectral/mandel/Mandel.hs, where the mandelset -- function gets a useful let-float if you inline windowToViewport -- NB: we used to have a second exception, for data con wrappers. -- On the grounds that we use gentle mode for rule LHSs, and -- they match better when data con wrappers are inlined. -- But that only really applies to the trivial wrappers (like (:)), -- and they are now constructed as Compulsory unfoldings (in MkId) -- so they'll happen anyway. SimplPhase n -> isActive n prag where prag = idInlinePragma id activeRule :: SimplEnv -> Maybe (Activation -> Bool) -- Nothing => No rules at all activeRule env | opt_RulesOff = Nothing | otherwise = case getMode env of SimplGently -> Just isAlwaysActive -- Used to be Nothing (no rules in gentle mode) -- Main motivation for changing is that I wanted -- lift String ===> ... -- to work in Template Haskell when simplifying -- splices, so we get simpler code for literal strings SimplPhase n -> Just (isActive n) \end{code} %************************************************************************ %* * Rebuilding a lambda %* * %************************************************************************ \begin{code} mkLam :: [OutBndr] -> OutExpr -> SimplM OutExpr -- mkLam tries three things -- a) eta reduction, if that gives a trivial expression -- b) eta expansion [only if there are some value lambdas] mkLam bndrs body = do { dflags <- getDOptsSmpl ; mkLam' dflags bndrs body } where mkLam' :: DynFlags -> [OutBndr] -> OutExpr -> SimplM OutExpr mkLam' dflags bndrs (Cast body@(Lam _ _) co) -- Note [Casts and lambdas] = do { lam <- mkLam' dflags (bndrs ++ bndrs') body' ; return (mkCoerce (mkPiTypes bndrs co) lam) } where (bndrs',body') = collectBinders body mkLam' dflags bndrs body | dopt Opt_DoEtaReduction dflags, Just etad_lam <- tryEtaReduce bndrs body = do { tick (EtaReduction (head bndrs)) ; return etad_lam } | dopt Opt_DoLambdaEtaExpansion dflags, any isRuntimeVar bndrs = do { body' <- tryEtaExpansion dflags body ; return (mkLams bndrs body') } | otherwise = returnSmpl (mkLams bndrs body) \end{code} Note [Casts and lambdas] ~~~~~~~~~~~~~~~~~~~~~~~~ Consider (\x. (\y. e) `cast` g1) `cast` g2 There is a danger here that the two lambdas look separated, and the full laziness pass might float an expression to between the two. So this equation in mkLam' floats the g1 out, thus: (\x. e `cast` g1) --> (\x.e) `cast` (tx -> g1) where x:tx. In general, this floats casts outside lambdas, where (I hope) they might meet and cancel with some other cast. -- c) floating lets out through big lambdas -- [only if all tyvar lambdas, and only if this lambda -- is the RHS of a let] {- Sept 01: I'm experimenting with getting the full laziness pass to float out past big lambdsa | all isTyVar bndrs, -- Only for big lambdas contIsRhs cont -- Only try the rhs type-lambda floating -- if this is indeed a right-hand side; otherwise -- we end up floating the thing out, only for float-in -- to float it right back in again! = tryRhsTyLam env bndrs body `thenSmpl` \ (floats, body') -> returnSmpl (floats, mkLams bndrs body') -} %************************************************************************ %* * \subsection{Eta expansion and reduction} %* * %************************************************************************ We try for eta reduction here, but *only* if we get all the way to an exprIsTrivial expression. We don't want to remove extra lambdas unless we are going to avoid allocating this thing altogether \begin{code} tryEtaReduce :: [OutBndr] -> OutExpr -> Maybe OutExpr tryEtaReduce bndrs body -- We don't use CoreUtils.etaReduce, because we can be more -- efficient here: -- (a) we already have the binders -- (b) we can do the triviality test before computing the free vars = go (reverse bndrs) body where go (b : bs) (App fun arg) | ok_arg b arg = go bs fun -- Loop round go [] fun | ok_fun fun = Just fun -- Success! go _ _ = Nothing -- Failure! ok_fun fun = exprIsTrivial fun && not (any (`elemVarSet` (exprFreeVars fun)) bndrs) && (exprIsHNF fun || all ok_lam bndrs) ok_lam v = isTyVar v || isDictId v -- The exprIsHNF is because eta reduction is not -- valid in general: \x. bot /= bot -- So we need to be sure that the "fun" is a value. -- -- However, we always want to reduce (/\a -> f a) to f -- This came up in a RULE: foldr (build (/\a -> g a)) -- did not match foldr (build (/\b -> ...something complex...)) -- The type checker can insert these eta-expanded versions, -- with both type and dictionary lambdas; hence the slightly -- ad-hoc isDictTy ok_arg b arg = varToCoreExpr b `cheapEqExpr` arg \end{code} Try eta expansion for RHSs We go for: f = \x1..xn -> N ==> f = \x1..xn y1..ym -> N y1..ym (n >= 0) where (in both cases) * The xi can include type variables * The yi are all value variables * N is a NORMAL FORM (i.e. no redexes anywhere) wanting a suitable number of extra args. We may have to sandwich some coerces between the lambdas to make the types work. exprEtaExpandArity looks through coerces when computing arity; and etaExpand adds the coerces as necessary when actually computing the expansion. \begin{code} tryEtaExpansion :: DynFlags -> OutExpr -> SimplM OutExpr -- There is at least one runtime binder in the binders tryEtaExpansion dflags body = getUniquesSmpl `thenSmpl` \ us -> returnSmpl (etaExpand fun_arity us body (exprType body)) where fun_arity = exprEtaExpandArity dflags body \end{code} %************************************************************************ %* * \subsection{Floating lets out of big lambdas} %* * %************************************************************************ tryRhsTyLam tries this transformation, when the big lambda appears as the RHS of a let(rec) binding: /\abc -> let(rec) x = e in b ==> let(rec) x' = /\abc -> let x = x' a b c in e in /\abc -> let x = x' a b c in b This is good because it can turn things like: let f = /\a -> letrec g = ... g ... in g into letrec g' = /\a -> ... g' a ... in let f = /\ a -> g' a which is better. In effect, it means that big lambdas don't impede let-floating. This optimisation is CRUCIAL in eliminating the junk introduced by desugaring mutually recursive definitions. Don't eliminate it lightly! So far as the implementation is concerned: Invariant: go F e = /\tvs -> F e Equalities: go F (Let x=e in b) = Let x' = /\tvs -> F e in go G b where G = F . Let x = x' tvs go F (Letrec xi=ei in b) = Letrec {xi' = /\tvs -> G ei} in go G b where G = F . Let {xi = xi' tvs} [May 1999] If we do this transformation *regardless* then we can end up with some pretty silly stuff. For example, let st = /\ s -> let { x1=r1 ; x2=r2 } in ... in .. becomes let y1 = /\s -> r1 y2 = /\s -> r2 st = /\s -> ...[y1 s/x1, y2 s/x2] in .. Unless the "..." is a WHNF there is really no point in doing this. Indeed it can make things worse. Suppose x1 is used strictly, and is of the form x1* = case f y of { (a,b) -> e } If we abstract this wrt the tyvar we then can't do the case inline as we would normally do. \begin{code} {- Trying to do this in full laziness tryRhsTyLam :: SimplEnv -> [OutTyVar] -> OutExpr -> SimplM FloatsWithExpr -- Call ensures that all the binders are type variables tryRhsTyLam env tyvars body -- Only does something if there's a let | not (all isTyVar tyvars) || not (worth_it body) -- inside a type lambda, = returnSmpl (emptyFloats env, body) -- and a WHNF inside that | otherwise = go env (\x -> x) body where worth_it e@(Let _ _) = whnf_in_middle e worth_it e = False whnf_in_middle (Let (NonRec x rhs) e) | isUnLiftedType (idType x) = False whnf_in_middle (Let _ e) = whnf_in_middle e whnf_in_middle e = exprIsCheap e main_tyvar_set = mkVarSet tyvars go env fn (Let bind@(NonRec var rhs) body) | exprIsTrivial rhs = go env (fn . Let bind) body go env fn (Let (NonRec var rhs) body) = mk_poly tyvars_here var `thenSmpl` \ (var', rhs') -> addAuxiliaryBind env (NonRec var' (mkLams tyvars_here (fn rhs))) $ \ env -> go env (fn . Let (mk_silly_bind var rhs')) body where tyvars_here = varSetElems (main_tyvar_set `intersectVarSet` exprSomeFreeVars isTyVar rhs) -- Abstract only over the type variables free in the rhs -- wrt which the new binding is abstracted. But the naive -- approach of abstract wrt the tyvars free in the Id's type -- fails. Consider: -- /\ a b -> let t :: (a,b) = (e1, e2) -- x :: a = fst t -- in ... -- Here, b isn't free in x's type, but we must nevertheless -- abstract wrt b as well, because t's type mentions b. -- Since t is floated too, we'd end up with the bogus: -- poly_t = /\ a b -> (e1, e2) -- poly_x = /\ a -> fst (poly_t a *b*) -- So for now we adopt the even more naive approach of -- abstracting wrt *all* the tyvars. We'll see if that -- gives rise to problems. SLPJ June 98 go env fn (Let (Rec prs) body) = mapAndUnzipSmpl (mk_poly tyvars_here) vars `thenSmpl` \ (vars', rhss') -> let gn body = fn (foldr Let body (zipWith mk_silly_bind vars rhss')) pairs = vars' `zip` [mkLams tyvars_here (gn rhs) | rhs <- rhss] in addAuxiliaryBind env (Rec pairs) $ \ env -> go env gn body where (vars,rhss) = unzip prs tyvars_here = varSetElems (main_tyvar_set `intersectVarSet` exprsSomeFreeVars isTyVar (map snd prs)) -- See notes with tyvars_here above go env fn body = returnSmpl (emptyFloats env, fn body) mk_poly tyvars_here var = getUniqueSmpl `thenSmpl` \ uniq -> let poly_name = setNameUnique (idName var) uniq -- Keep same name poly_ty = mkForAllTys tyvars_here (idType var) -- But new type of course poly_id = mkLocalId poly_name poly_ty -- In the olden days, it was crucial to copy the occInfo of the original var, -- because we were looking at occurrence-analysed but as yet unsimplified code! -- In particular, we mustn't lose the loop breakers. BUT NOW we are looking -- at already simplified code, so it doesn't matter -- -- It's even right to retain single-occurrence or dead-var info: -- Suppose we started with /\a -> let x = E in B -- where x occurs once in B. Then we transform to: -- let x' = /\a -> E in /\a -> let x* = x' a in B -- where x* has an INLINE prag on it. Now, once x* is inlined, -- the occurrences of x' will be just the occurrences originally -- pinned on x. in returnSmpl (poly_id, mkTyApps (Var poly_id) (mkTyVarTys tyvars_here)) mk_silly_bind var rhs = NonRec var (Note InlineMe rhs) -- Suppose we start with: -- -- x = /\ a -> let g = G in E -- -- Then we'll float to get -- -- x = let poly_g = /\ a -> G -- in /\ a -> let g = poly_g a in E -- -- But now the occurrence analyser will see just one occurrence -- of poly_g, not inside a lambda, so the simplifier will -- PreInlineUnconditionally poly_g back into g! Badk to square 1! -- (I used to think that the "don't inline lone occurrences" stuff -- would stop this happening, but since it's the *only* occurrence, -- PreInlineUnconditionally kicks in first!) -- -- Solution: put an INLINE note on g's RHS, so that poly_g seems -- to appear many times. (NB: mkInlineMe eliminates -- such notes on trivial RHSs, so do it manually.) -} \end{code} %************************************************************************ %* * prepareAlts %* * %************************************************************************ prepareAlts tries these things: 1. If several alternatives are identical, merge them into a single DEFAULT alternative. I've occasionally seen this making a big difference: case e of =====> case e of C _ -> f x D v -> ....v.... D v -> ....v.... DEFAULT -> f x DEFAULT -> f x The point is that we merge common RHSs, at least for the DEFAULT case. [One could do something more elaborate but I've never seen it needed.] To avoid an expensive test, we just merge branches equal to the *first* alternative; this picks up the common cases a) all branches equal b) some branches equal to the DEFAULT (which occurs first) 2. Case merging: case e of b { ==> case e of b { p1 -> rhs1 p1 -> rhs1 ... ... pm -> rhsm pm -> rhsm _ -> case b of b' { pn -> let b'=b in rhsn pn -> rhsn ... ... po -> let b'=b in rhso po -> rhso _ -> let b'=b in rhsd _ -> rhsd } which merges two cases in one case when -- the default alternative of the outer case scrutises the same variable as the outer case This transformation is called Case Merging. It avoids that the same variable is scrutinised multiple times. The case where transformation (1) showed up was like this (lib/std/PrelCError.lhs): x | p `is` 1 -> e1 | p `is` 2 -> e2 ...etc... where @is@ was something like p `is` n = p /= (-1) && p == n This gave rise to a horrible sequence of cases case p of (-1) -> $j p 1 -> e1 DEFAULT -> $j p and similarly in cascade for all the join points! Note [Dead binders] ~~~~~~~~~~~~~~~~~~~~ We do this *here*, looking at un-simplified alternatives, because we have to check that r doesn't mention the variables bound by the pattern in each alternative, so the binder-info is rather useful. \begin{code} prepareAlts :: OutExpr -> OutId -> [InAlt] -> SimplM ([AltCon], [InAlt]) prepareAlts scrut case_bndr' alts = do { dflags <- getDOptsSmpl ; alts <- combineIdenticalAlts case_bndr' alts ; let (alts_wo_default, maybe_deflt) = findDefault alts alt_cons = [con | (con,_,_) <- alts_wo_default] imposs_deflt_cons = nub (imposs_cons ++ alt_cons) -- "imposs_deflt_cons" are handled -- EITHER by the context, -- OR by a non-DEFAULT branch in this case expression. ; default_alts <- prepareDefault dflags scrut case_bndr' mb_tc_app imposs_deflt_cons maybe_deflt ; let trimmed_alts = filter possible_alt alts_wo_default merged_alts = mergeAlts trimmed_alts default_alts -- We need the mergeAlts in case the new default_alt -- has turned into a constructor alternative. -- The merge keeps the inner DEFAULT at the front, if there is one -- and interleaves the alternatives in the right order ; return (imposs_deflt_cons, merged_alts) } where mb_tc_app = splitTyConApp_maybe (idType case_bndr') Just (_, inst_tys) = mb_tc_app imposs_cons = case scrut of Var v -> otherCons (idUnfolding v) other -> [] possible_alt :: CoreAlt -> Bool possible_alt (con, _, _) | con `elem` imposs_cons = False possible_alt (DataAlt con, _, _) = dataConCanMatch inst_tys con possible_alt alt = True -------------------------------------------------- -- 1. Merge identical branches -------------------------------------------------- combineIdenticalAlts :: OutId -> [InAlt] -> SimplM [InAlt] combineIdenticalAlts case_bndr alts@((con1,bndrs1,rhs1) : con_alts) | all isDeadBinder bndrs1, -- Remember the default length filtered_alts < length con_alts -- alternative comes first -- Also Note [Dead binders] = do { tick (AltMerge case_bndr) ; return ((DEFAULT, [], rhs1) : filtered_alts) } where filtered_alts = filter keep con_alts keep (con,bndrs,rhs) = not (all isDeadBinder bndrs && rhs `cheapEqExpr` rhs1) combineIdenticalAlts case_bndr alts = return alts ------------------------------------------------------------------------- -- Prepare the default alternative ------------------------------------------------------------------------- prepareDefault :: DynFlags -> OutExpr -- Scrutinee -> OutId -- Case binder; need just for its type. Note that as an -- OutId, it has maximum information; this is important. -- Test simpl013 is an example -> Maybe (TyCon, [Type]) -- Type of scrutinee, decomposed -> [AltCon] -- These cons can't happen when matching the default -> Maybe InExpr -- Rhs -> SimplM [InAlt] -- Still unsimplified -- We use a list because it's what mergeAlts expects, -- And becuase case-merging can cause many to show up ------- Merge nested cases ---------- prepareDefault dflags scrut outer_bndr bndr_ty imposs_cons (Just deflt_rhs) | dopt Opt_CaseMerge dflags , Case (Var scrut_var) inner_bndr _ inner_alts <- deflt_rhs , scruting_same_var scrut_var = do { tick (CaseMerge outer_bndr) ; let munge_rhs rhs = bindCaseBndr inner_bndr (Var outer_bndr) rhs ; return [(con, args, munge_rhs rhs) | (con, args, rhs) <- inner_alts, not (con `elem` imposs_cons) ] -- NB: filter out any imposs_cons. Example: -- case x of -- A -> e1 -- DEFAULT -> case x of -- A -> e2 -- B -> e3 -- When we merge, we must ensure that e1 takes -- precedence over e2 as the value for A! } -- Warning: don't call prepareAlts recursively! -- Firstly, there's no point, because inner alts have already had -- mkCase applied to them, so they won't have a case in their default -- Secondly, if you do, you get an infinite loop, because the bindCaseBndr -- in munge_rhs may put a case into the DEFAULT branch! where -- We are scrutinising the same variable if it's -- the outer case-binder, or if the outer case scrutinises a variable -- (and it's the same). Testing both allows us not to replace the -- outer scrut-var with the outer case-binder (Simplify.simplCaseBinder). scruting_same_var = case scrut of Var outer_scrut -> \ v -> v == outer_bndr || v == outer_scrut other -> \ v -> v == outer_bndr --------- Fill in known constructor ----------- prepareDefault dflags scrut case_bndr (Just (tycon, inst_tys)) imposs_cons (Just deflt_rhs) | -- This branch handles the case where we are -- scrutinisng an algebraic data type isAlgTyCon tycon -- It's a data type, tuple, or unboxed tuples. , not (isNewTyCon tycon) -- We can have a newtype, if we are just doing an eval: -- case x of { DEFAULT -> e } -- and we don't want to fill in a default for them! , Just all_cons <- tyConDataCons_maybe tycon , not (null all_cons) -- This is a tricky corner case. If the data type has no constructors, -- which GHC allows, then the case expression will have at most a default -- alternative. We don't want to eliminate that alternative, because the -- invariant is that there's always one alternative. It's more convenient -- to leave -- case x of { DEFAULT -> e } -- as it is, rather than transform it to -- error "case cant match" -- which would be quite legitmate. But it's a really obscure corner, and -- not worth wasting code on. , let imposs_data_cons = [con | DataAlt con <- imposs_cons] -- We now know it's a data type is_possible con = not (con `elem` imposs_data_cons) && dataConCanMatch inst_tys con = case filter is_possible all_cons of [] -> return [] -- Eliminate the default alternative -- altogether if it can't match [con] -> -- It matches exactly one constructor, so fill it in do { tick (FillInCaseDefault case_bndr) ; us <- getUniquesSmpl ; let (ex_tvs, co_tvs, arg_ids) = dataConRepInstPat us con inst_tys ; return [(DataAlt con, ex_tvs ++ co_tvs ++ arg_ids, deflt_rhs)] } two_or_more -> return [(DEFAULT, [], deflt_rhs)] --------- Catch-all cases ----------- prepareDefault dflags scrut case_bndr bndr_ty imposs_cons (Just deflt_rhs) = return [(DEFAULT, [], deflt_rhs)] prepareDefault dflags scrut case_bndr bndr_ty imposs_cons Nothing = return [] -- No default branch \end{code} ================================================================================= mkCase tries these things 1. Eliminate the case altogether if possible 2. Case-identity: case e of ===> e True -> True; False -> False and similar friends. \begin{code} mkCase :: OutExpr -> OutId -> OutType -> [OutAlt] -- Increasing order -> SimplM OutExpr -------------------------------------------------- -- 1. Check for empty alternatives -------------------------------------------------- -- This isn't strictly an error. It's possible that the simplifer might "see" -- that an inner case has no accessible alternatives before it "sees" that the -- entire branch of an outer case is inaccessible. So we simply -- put an error case here insteadd mkCase scrut case_bndr ty [] = pprTrace "mkCase: null alts" (ppr case_bndr <+> ppr scrut) $ return (mkApps (Var eRROR_ID) [Type ty, Lit (mkStringLit "Impossible alternative")]) -------------------------------------------------- -- 2. Identity case -------------------------------------------------- mkCase scrut case_bndr ty alts -- Identity case | all identity_alt alts = tick (CaseIdentity case_bndr) `thenSmpl_` returnSmpl (re_cast scrut) where identity_alt (con, args, rhs) = check_eq con args (de_cast rhs) check_eq DEFAULT _ (Var v) = v == case_bndr check_eq (LitAlt lit') _ (Lit lit) = lit == lit' check_eq (DataAlt con) args rhs = rhs `cheapEqExpr` mkConApp con (arg_tys ++ varsToCoreExprs args) || rhs `cheapEqExpr` Var case_bndr check_eq con args rhs = False arg_tys = map Type (tyConAppArgs (idType case_bndr)) -- We've seen this: -- case e of x { _ -> x `cast` c } -- And we definitely want to eliminate this case, to give -- e `cast` c -- So we throw away the cast from the RHS, and reconstruct -- it at the other end. All the RHS casts must be the same -- if (all identity_alt alts) holds. -- -- Don't worry about nested casts, because the simplifier combines them de_cast (Cast e _) = e de_cast e = e re_cast scrut = case head alts of (_,_,Cast _ co) -> Cast scrut co other -> scrut -------------------------------------------------- -- Catch-all -------------------------------------------------- mkCase scrut bndr ty alts = returnSmpl (Case scrut bndr ty alts) \end{code} When adding auxiliary bindings for the case binder, it's worth checking if its dead, because it often is, and occasionally these mkCase transformations cascade rather nicely. \begin{code} bindCaseBndr bndr rhs body | isDeadBinder bndr = body | otherwise = bindNonRec bndr rhs body \end{code}