{-# OPTIONS -fno-warn-tabs #-} module DmdAnal ( dmdAnalProgram ) where #include "HsVersions.h" import Var ( isTyVar ) import DynFlags import WwLib ( deepSplitProductType_maybe ) import Demand -- All of it import CoreSyn import Outputable import VarEnv import BasicTypes import FastString import Data.List import DataCon import Id import CoreUtils ( exprIsHNF, exprType, exprIsTrivial ) -- import PprCore import TyCon import Type ( eqType ) -- import Pair -- import Coercion ( coercionKind ) import FamInstEnv import Util import Maybes ( isJust ) import TysWiredIn ( unboxedPairDataCon ) import TysPrim ( realWorldStatePrimTy ) import ErrUtils ( dumpIfSet_dyn ) import Name ( getName, stableNameCmp ) import Data.Function ( on )\end{code} %************************************************************************ %* * \subsection{Top level stuff} %* * %************************************************************************ \begin{code}

dmdAnalProgram :: DynFlags -> FamInstEnvs -> CoreProgram -> IO CoreProgram dmdAnalProgram dflags fam_envs binds = do { let { binds_plus_dmds = do_prog binds } ; dumpIfSet_dyn dflags Opt_D_dump_strsigs "Strictness signatures" $ dumpStrSig binds_plus_dmds ; return binds_plus_dmds } where do_prog :: CoreProgram -> CoreProgram do_prog binds = snd $ mapAccumL dmdAnalTopBind (emptyAnalEnv dflags fam_envs) binds -- Analyse a (group of) top-level binding(s) dmdAnalTopBind :: AnalEnv -> CoreBind -> (AnalEnv, CoreBind) dmdAnalTopBind sigs (NonRec id rhs) = (extendAnalEnv TopLevel sigs id sig, NonRec id2 rhs2) where ( _, _, _, rhs1) = dmdAnalRhs TopLevel Nothing sigs id rhs (sig, _, id2, rhs2) = dmdAnalRhs TopLevel Nothing (nonVirgin sigs) id rhs1 -- Do two passes to improve CPR information -- See comments with ignore_cpr_info in mk_sig_ty -- and with extendSigsWithLam dmdAnalTopBind sigs (Rec pairs) = (sigs', Rec pairs') where (sigs', _, pairs') = dmdFix TopLevel sigs pairs -- We get two iterations automatically -- c.f. the NonRec case above\end{code} %************************************************************************ %* * \subsection{The analyser itself} %* * %************************************************************************ Note [Ensure demand is strict] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ It's important not to analyse e with a lazy demand because a) When we encounter case s of (a,b) -> we demand s with U(d1d2)... but if the overall demand is lazy that is wrong, and we'd need to reduce the demand on s, which is inconvenient b) More important, consider f (let x = R in x+x), where f is lazy We still want to mark x as demanded, because it will be when we enter the let. If we analyse f's arg with a Lazy demand, we'll just mark x as Lazy c) The application rule wouldn't be right either Evaluating (f x) in a L demand does *not* cause evaluation of f in a C(L) demand! \begin{code}

-- If e is complicated enough to become a thunk, its contents will be evaluated -- at most once, so oneify it. dmdTransformThunkDmd :: CoreExpr -> Demand -> Demand dmdTransformThunkDmd e | exprIsTrivial e = id | otherwise = oneifyDmd -- Do not process absent demands -- Otherwise act like in a normal demand analysis -- See |-* relation in the companion paper dmdAnalStar :: AnalEnv -> Demand -- This one takes a *Demand* -> CoreExpr -> (BothDmdArg, CoreExpr) dmdAnalStar env dmd e | (cd, defer_and_use) <- toCleanDmd dmd (exprType e) , (dmd_ty, e') <- dmdAnal env cd e = (postProcessDmdTypeM defer_and_use dmd_ty, e') -- Main Demand Analsysis machinery dmdAnal :: AnalEnv -> CleanDemand -- The main one takes a *CleanDemand* -> CoreExpr -> (DmdType, CoreExpr) -- The CleanDemand is always strict and not absent -- See Note [Ensure demand is strict] dmdAnal _ _ (Lit lit) = (nopDmdType, Lit lit) dmdAnal _ _ (Type ty) = (nopDmdType, Type ty) -- Doesn't happen, in fact dmdAnal _ _ (Coercion co) = (nopDmdType, Coercion co) dmdAnal env dmd (Var var) = (dmdTransform env var dmd, Var var) dmdAnal env dmd (Cast e co) = (dmd_ty, Cast e' co) where (dmd_ty, e') = dmdAnal env dmd e {- ----- I don't get this, so commenting out ------- to_co = pSnd (coercionKind co) dmd' | Just tc <- tyConAppTyCon_maybe to_co , isRecursiveTyCon tc = cleanEvalDmd | otherwise = dmd -- This coerce usually arises from a recursive -- newtype, and we don't want to look inside them -- for exactly the same reason that we don't look -- inside recursive products -- we might not reach -- a fixpoint. So revert to a vanilla Eval demand -} dmdAnal env dmd (Tick t e) = (dmd_ty, Tick t e') where (dmd_ty, e') = dmdAnal env dmd e dmdAnal env dmd (App fun (Type ty)) = (fun_ty, App fun' (Type ty)) where (fun_ty, fun') = dmdAnal env dmd fun dmdAnal sigs dmd (App fun (Coercion co)) = (fun_ty, App fun' (Coercion co)) where (fun_ty, fun') = dmdAnal sigs dmd fun -- Lots of the other code is there to make this -- beautiful, compositional, application rule :-) dmdAnal env dmd (App fun arg) -- Non-type arguments = let -- [Type arg handled above] call_dmd = mkCallDmd dmd (fun_ty, fun') = dmdAnal env call_dmd fun (arg_dmd, res_ty) = splitDmdTy fun_ty (arg_ty, arg') = dmdAnalStar env (dmdTransformThunkDmd arg arg_dmd) arg in -- pprTrace "dmdAnal:app" (vcat -- [ text "dmd =" <+> ppr dmd -- , text "expr =" <+> ppr (App fun arg) -- , text "fun dmd_ty =" <+> ppr fun_ty -- , text "arg dmd =" <+> ppr arg_dmd -- , text "arg dmd_ty =" <+> ppr arg_ty -- , text "res dmd_ty =" <+> ppr res_ty -- , text "overall res dmd_ty =" <+> ppr (res_ty `bothDmdType` arg_ty) ]) (res_ty `bothDmdType` arg_ty, App fun' arg') -- this is an anonymous lambda, since @dmdAnalRhs@ uses @collectBinders@ dmdAnal env dmd (Lam var body) | isTyVar var = let (body_ty, body') = dmdAnal env dmd body in (body_ty, Lam var body') | otherwise = let (body_dmd, defer_and_use@(_,one_shot)) = peelCallDmd dmd -- body_dmd - a demand to analyze the body -- one_shot - one-shotness of the lambda -- hence, cardinality of its free vars env' = extendSigsWithLam env var (body_ty, body') = dmdAnal env' body_dmd body (lam_ty, var') = annotateLamIdBndr env notArgOfDfun body_ty one_shot var in (postProcessUnsat defer_and_use lam_ty, Lam var' body') dmdAnal env dmd (Case scrut case_bndr ty [alt@(DataAlt dc, _, _)]) -- Only one alternative with a product constructor | let tycon = dataConTyCon dc , isProductTyCon tycon , Just rec_tc' <- checkRecTc (ae_rec_tc env) tycon = let env_w_tc = env { ae_rec_tc = rec_tc' } env_alt = extendAnalEnv NotTopLevel env_w_tc case_bndr case_bndr_sig (alt_ty, alt') = dmdAnalAlt env_alt dmd alt (alt_ty1, case_bndr') = annotateBndr env alt_ty case_bndr (_, bndrs', _) = alt' case_bndr_sig = cprProdSig (dataConRepArity dc) -- Inside the alternative, the case binder has the CPR property. -- Meaning that a case on it will successfully cancel. -- Example: -- f True x = case x of y { I# x' -> if x' ==# 3 then y else I# 8 } -- f False x = I# 3 -- -- We want f to have the CPR property: -- f b x = case fw b x of { r -> I# r } -- fw True x = case x of y { I# x' -> if x' ==# 3 then x' else 8 } -- fw False x = 3 -- Figure out whether the demand on the case binder is used, and use -- that to set the scrut_dmd. This is utterly essential. -- Consider f x = case x of y { (a,b) -> k y a } -- If we just take scrut_demand = U(L,A), then we won't pass x to the -- worker, so the worker will rebuild -- x = (a, absent-error) -- and that'll crash. -- So at one stage I had: -- dead_case_bndr = isAbsDmd (idDemandInfo case_bndr') -- keepity | dead_case_bndr = Drop -- | otherwise = Keep -- -- But then consider -- case x of y { (a,b) -> h y + a } -- where h : U(LL) -> T -- The above code would compute a Keep for x, since y is not Abs, which is silly -- The insight is, of course, that a demand on y is a demand on the -- scrutinee, so we need to `both` it with the scrut demand scrut_dmd1 = mkProdDmd [idDemandInfo b | b <- bndrs', isId b] scrut_dmd2 = strictenDmd (idDemandInfo case_bndr') scrut_dmd = scrut_dmd1 `bothCleanDmd` scrut_dmd2 (scrut_ty, scrut') = dmdAnal env scrut_dmd scrut res_ty = alt_ty1 `bothDmdType` toBothDmdArg scrut_ty in -- pprTrace "dmdAnal:Case1" (vcat [ text "scrut" <+> ppr scrut -- , text "dmd" <+> ppr dmd -- , text "case_bndr_dmd" <+> ppr (idDemandInfo case_bndr') -- , text "scrut_dmd" <+> ppr scrut_dmd -- , text "scrut_ty" <+> ppr scrut_ty -- , text "alt_ty" <+> ppr alt_ty1 -- , text "res_ty" <+> ppr res_ty ]) $ (res_ty, Case scrut' case_bndr' ty [alt']) dmdAnal env dmd (Case scrut case_bndr ty alts) = let -- Case expression with multiple alternatives (alt_tys, alts') = mapAndUnzip (dmdAnalAlt env dmd) alts (scrut_ty, scrut') = dmdAnal env cleanEvalDmd scrut (alt_ty, case_bndr') = annotateBndr env (foldr lubDmdType botDmdType alt_tys) case_bndr res_ty = alt_ty `bothDmdType` toBothDmdArg scrut_ty in -- pprTrace "dmdAnal:Case2" (vcat [ text "scrut" <+> ppr scrut -- , text "scrut_ty" <+> ppr scrut_ty -- , text "alt_tys" <+> ppr alt_tys -- , text "alt_ty" <+> ppr alt_ty -- , text "res_ty" <+> ppr res_ty ]) $ (res_ty, Case scrut' case_bndr' ty alts') dmdAnal env dmd (Let (NonRec id rhs) body) = (body_ty2, Let (NonRec id2 annotated_rhs) body') where (sig, lazy_fv, id1, rhs') = dmdAnalRhs NotTopLevel Nothing env id rhs (body_ty, body') = dmdAnal (extendAnalEnv NotTopLevel env id sig) dmd body (body_ty1, id2) = annotateBndr env body_ty id1 body_ty2 = addLazyFVs body_ty1 lazy_fv -- Annotate top-level lambdas at RHS basing on the aggregated demand info -- See Note [Annotating lambdas at right-hand side] annotated_rhs = annLamWithShotness (idDemandInfo id2) rhs' -- If the actual demand is better than the vanilla call -- demand, you might think that we might do better to re-analyse -- the RHS with the stronger demand. -- But (a) That seldom happens, because it means that *every* path in -- the body of the let has to use that stronger demand -- (b) It often happens temporarily in when fixpointing, because -- the recursive function at first seems to place a massive demand. -- But we don't want to go to extra work when the function will -- probably iterate to something less demanding. -- In practice, all the times the actual demand on id2 is more than -- the vanilla call demand seem to be due to (b). So we don't -- bother to re-analyse the RHS. dmdAnal env dmd (Let (Rec pairs) body) = let (env', lazy_fv, pairs') = dmdFix NotTopLevel env pairs (body_ty, body') = dmdAnal env' dmd body body_ty1 = deleteFVs body_ty (map fst pairs) body_ty2 = addLazyFVs body_ty1 lazy_fv in body_ty2 `seq` (body_ty2, Let (Rec pairs') body') annLamWithShotness :: Demand -> CoreExpr -> CoreExpr annLamWithShotness d e | Just u <- cleanUseDmd_maybe d = go u e | otherwise = e where go u e | Just (c, u') <- peelUseCall u , Lam bndr body <- e = if isTyVar bndr then Lam bndr (go u body) else Lam (setOneShotness c bndr) (go u' body) | otherwise = e setOneShotness :: Count -> Id -> Id setOneShotness One bndr = setOneShotLambda bndr setOneShotness Many bndr = bndr dmdAnalAlt :: AnalEnv -> CleanDemand -> Alt Var -> (DmdType, Alt Var) dmdAnalAlt env dmd (con,bndrs,rhs) = let (rhs_ty, rhs') = dmdAnal env dmd rhs rhs_ty' = addDataConPatDmds con bndrs rhs_ty (alt_ty, bndrs') = annotateBndrs env rhs_ty' bndrs final_alt_ty | io_hack_reqd = deferAfterIO alt_ty | otherwise = alt_ty -- Note [IO hack in the demand analyser] -- -- There's a hack here for I/O operations. Consider -- case foo x s of { (# s, r #) -> y } -- Is this strict in 'y'. Normally yes, but what if 'foo' is an I/O -- operation that simply terminates the program (not in an erroneous way)? -- In that case we should not evaluate y before the call to 'foo'. -- Hackish solution: spot the IO-like situation and add a virtual branch, -- as if we had -- case foo x s of -- (# s, r #) -> y -- other -> return () -- So the 'y' isn't necessarily going to be evaluated -- -- A more complete example (Trac #148, #1592) where this shows up is: -- do { let len = <expensive> ; -- ; when (...) (exitWith ExitSuccess) -- ; print len } io_hack_reqd = con == DataAlt unboxedPairDataCon && idType (head bndrs) `eqType` realWorldStatePrimTy in (final_alt_ty, (con, bndrs', rhs'))\end{code} Note [Aggregated demand for cardinality] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We use different strategies for strictness and usage/cardinality to "unleash" demands captured on free variables by bindings. Let us consider the example: f1 y = let {-# NOINLINE h #-} h = y in (h, h) We are interested in obtaining cardinality demand U1 on |y|, as it is used only in a thunk, and, therefore, is not going to be updated any more. Therefore, the demand on |y|, captured and unleashed by usage of |h| is U1. However, if we unleash this demand every time |h| is used, and then sum up the effects, the ultimate demand on |y| will be U1 + U1 = U. In order to avoid it, we *first* collect the aggregate demand on |h| in the body of let-expression, and only then apply the demand transformer: transf[x](U) = {y |-> U1} so the resulting demand on |y| is U1. The situation is, however, different for strictness, where this aggregating approach exhibits worse results because of the nature of |both| operation for strictness. Consider the example: f y c = let h x = y |seq| x in case of True -> h True False -> y It is clear that |f| is strict in |y|, however, the suggested analysis will infer from the body of |let| that |h| is used lazily (as it is used in one branch only), therefore lazy demand will be put on its free variable |y|. Conversely, if the demand on |h| is unleashed right on the spot, we will get the desired result, namely, that |f| is strict in |y|. Note [Annotating lambdas at right-hand side] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Let us take a look at the following example: g f = let x = 100 h = \y -> f x y in h 5 One can see that |h| is called just once, therefore the RHS of h can be annotated as a one-shot lambda. This is done by the function annLamWithShotness *a posteriori*, i.e., basing on the aggregated usage demand on |h| from the body of |let|-expression, which is C1(U) in this case. In other words, for locally-bound lambdas we can infer one-shotness. \begin{code}

addDataConPatDmds :: AltCon -> [Var] -> DmdType -> DmdType -- See Note [Add demands for strict constructors] addDataConPatDmds DEFAULT _ dmd_ty = dmd_ty addDataConPatDmds (LitAlt _) _ dmd_ty = dmd_ty addDataConPatDmds (DataAlt con) bndrs dmd_ty = foldr add dmd_ty str_bndrs where add bndr dmd_ty = addVarDmd dmd_ty bndr seqDmd str_bndrs = [ b | (b,s) <- zipEqual "addDataConPatBndrs" (filter isId bndrs) (dataConRepStrictness con) , isMarkedStrict s ]\end{code} Note [Add demands for strict constructors] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider this program (due to Roman): data X a = X !a foo :: X Int -> Int -> Int foo (X a) n = go 0 where go i | i < n = a + go (i+1) | otherwise = 0 We want the worker for 'foo' too look like this: $wfoo :: Int# -> Int# -> Int# with the first argument unboxed, so that it is not eval'd each time around the loop (which would otherwise happen, since 'foo' is not strict in 'a'. It is sound for the wrapper to pass an unboxed arg because X is strict, so its argument must be evaluated. And if we *don't* pass an unboxed argument, we can't even repair it by adding a `seq` thus: foo (X a) n = a `seq` go 0 because the seq is discarded (very early) since X is strict! There is the usual danger of reboxing, which as usual we ignore. But if X is monomorphic, and has an UNPACK pragma, then this optimisation is even more important. We don't want the wrapper to rebox an unboxed argument, and pass an Int to $wfoo! %************************************************************************ %* * Demand transformer %* * %************************************************************************ \begin{code}

dmdTransform :: AnalEnv -- The strictness environment -> Id -- The function -> CleanDemand -- The demand on the function -> DmdType -- The demand type of the function in this context -- Returned DmdEnv includes the demand on -- this function plus demand on its free variables dmdTransform env var dmd | isDataConWorkId var -- Data constructor = dmdTransformDataConSig (idArity var) (idStrictness var) dmd | gopt Opt_DmdTxDictSel (ae_dflags env), Just _ <- isClassOpId_maybe var -- Dictionary component selector = dmdTransformDictSelSig (idStrictness var) dmd | isGlobalId var -- Imported function = let res = dmdTransformSig (idStrictness var) dmd in -- pprTrace "dmdTransform" (vcat [ppr var, ppr (idStrictness var), ppr dmd, ppr res]) res | Just (sig, top_lvl) <- lookupSigEnv env var -- Local letrec bound thing , let fn_ty = dmdTransformSig sig dmd = -- pprTrace "dmdTransform" (vcat [ppr var, ppr sig, ppr dmd, ppr fn_ty]) $ if isTopLevel top_lvl then fn_ty -- Don't record top level things else addVarDmd fn_ty var (mkOnceUsedDmd dmd) | otherwise -- Local non-letrec-bound thing = unitVarDmd var (mkOnceUsedDmd dmd)\end{code} %************************************************************************ %* * \subsection{Bindings} %* * %************************************************************************ \begin{code}

-- Recursive bindings dmdFix :: TopLevelFlag -> AnalEnv -- Does not include bindings for this binding -> [(Id,CoreExpr)] -> (AnalEnv, DmdEnv, [(Id,CoreExpr)]) -- Binders annotated with stricness info dmdFix top_lvl env orig_pairs = (updSigEnv env (sigEnv final_env), lazy_fv, pairs') -- Return to original virgin state, keeping new signatures where bndrs = map fst orig_pairs initial_env = addInitialSigs top_lvl env bndrs (final_env, lazy_fv, pairs') = loop 1 initial_env orig_pairs loop :: Int -> AnalEnv -- Already contains the current sigs -> [(Id,CoreExpr)] -> (AnalEnv, DmdEnv, [(Id,CoreExpr)]) loop n env pairs = -- pprTrace "dmd loop" (ppr n <+> ppr bndrs $$ ppr env) $ loop' n env pairs loop' n env pairs | found_fixpoint = (env', lazy_fv, pairs') -- Note: return pairs', not pairs. pairs' is the result of -- processing the RHSs with sigs (= sigs'), whereas pairs -- is the result of processing the RHSs with the *previous* -- iteration of sigs. | n >= 10 = -- pprTrace "dmdFix loop" (ppr n <+> (vcat -- [ text "Sigs:" <+> ppr [ (id,lookupVarEnv (sigEnv env) id, -- lookupVarEnv (sigEnv env') id) -- | (id,_) <- pairs], -- text "env:" <+> ppr env, -- text "binds:" <+> pprCoreBinding (Rec pairs)])) (env, lazy_fv, orig_pairs) -- Safe output -- The lazy_fv part is really important! orig_pairs has no strictness -- info, including nothing about free vars. But if we have -- letrec f = ....y..... in ...f... -- where 'y' is free in f, we must record that y is mentioned, -- otherwise y will get recorded as absent altogether | otherwise = loop (n+1) (nonVirgin env') pairs' where found_fixpoint = all (same_sig (sigEnv env) (sigEnv env')) bndrs ((env',lazy_fv), pairs') = mapAccumL my_downRhs (env, emptyDmdEnv) pairs -- mapAccumL: Use the new signature to do the next pair -- The occurrence analyser has arranged them in a good order -- so this can significantly reduce the number of iterations needed my_downRhs (env, lazy_fv) (id,rhs) = ((env', lazy_fv'), (id', rhs')) where (sig, lazy_fv1, id', rhs') = dmdAnalRhs top_lvl (Just bndrs) env id rhs lazy_fv' = plusVarEnv_C bothDmd lazy_fv lazy_fv1 env' = extendAnalEnv top_lvl env id sig same_sig sigs sigs' var = lookup sigs var == lookup sigs' var lookup sigs var = case lookupVarEnv sigs var of Just (sig,_) -> sig Nothing -> pprPanic "dmdFix" (ppr var) -- Non-recursive bindings dmdAnalRhs :: TopLevelFlag -> Maybe [Id] -- Just bs <=> recursive, Nothing <=> non-recursive -> AnalEnv -> Id -> CoreExpr -> (StrictSig, DmdEnv, Id, CoreExpr) -- Process the RHS of the binding, add the strictness signature -- to the Id, and augment the environment with the signature as well. dmdAnalRhs top_lvl rec_flag env id rhs | Just fn <- unpackTrivial rhs -- See Note [Demand analysis for trivial right-hand sides] , let fn_str = getStrictness env fn fn_fv | isLocalId fn = unitVarEnv fn topDmd | otherwise = emptyDmdEnv -- Note [Remember to demand the function itself] -- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ -- fn_fv: don't forget to produce a demand for fn itself -- Lacking this caused Trac #9128 -- The demand is very conservative (topDmd), but that doesn't -- matter; trivial bindings are usually inlined, so it only -- kicks in for top-level bindings and NOINLINE bindings = (fn_str, fn_fv, set_idStrictness env id fn_str, rhs) | otherwise = (sig_ty, lazy_fv, id', mkLams bndrs' body') where (bndrs, body) = collectBinders rhs env_body = foldl extendSigsWithLam env bndrs (body_ty, body') = dmdAnal env_body body_dmd body body_ty' = removeDmdTyArgs body_ty -- zap possible deep CPR info (DmdType rhs_fv rhs_dmds rhs_res, bndrs') = annotateLamBndrs env (isDFunId id) body_ty' bndrs sig_ty = mkStrictSig (mkDmdType sig_fv rhs_dmds rhs_res') id' = set_idStrictness env id sig_ty -- See Note [NOINLINE and strictness] -- See Note [Product demands for function body] body_dmd = case deepSplitProductType_maybe (ae_fam_envs env) (exprType body) of Nothing -> cleanEvalDmd Just (dc, _, _, _) -> cleanEvalProdDmd (dataConRepArity dc) -- See Note [Lazy and unleashable free variables] -- See Note [Aggregated demand for cardinality] rhs_fv1 = case rec_flag of Just bs -> reuseEnv (delVarEnvList rhs_fv bs) Nothing -> rhs_fv (lazy_fv, sig_fv) = splitFVs is_thunk rhs_fv1 rhs_res' = trimCPRInfo trim_all trim_sums rhs_res trim_all = is_thunk && not_strict trim_sums = not (isTopLevel top_lvl) -- See Note [CPR for sum types] -- See Note [CPR for thunks] is_thunk = not (exprIsHNF rhs) not_strict = isTopLevel top_lvl -- Top level and recursive things don't || isJust rec_flag -- get their demandInfo set at all || not (isStrictDmd (idDemandInfo id) || ae_virgin env) -- See Note [Optimistic CPR in the "virgin" case] unpackTrivial :: CoreExpr -> Maybe Id -- Returns (Just v) if the arg is really equal to v, modulo -- casts, type applications etc -- See Note [Demand analysis for trivial right-hand sides] unpackTrivial (Var v) = Just v unpackTrivial (Cast e _) = unpackTrivial e unpackTrivial (Lam v e) | isTyVar v = unpackTrivial e unpackTrivial (App e a) | isTypeArg a = unpackTrivial e unpackTrivial _ = Nothing\end{code} Note [Demand analysis for trivial right-hand sides] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider foo = plusInt |> co where plusInt is an arity-2 function with known strictness. Clearly we want plusInt's strictness to propagate to foo! But because it has no manifest lambdas, it won't do so automatically, and indeed 'co' might have type (Int->Int->Int) ~ T, so we *can't* eta-expand. So we have a special case for right-hand sides that are "trivial", namely variables, casts, type applications, and the like. Note that this can mean that 'foo' has an arity that is smaller than that indicated by its demand info. e.g. if co :: (Int->Int->Int) ~ T, then foo's arity will be zero (see Note [exprArity invariant] in CoreArity), but its demand signature will be that of plusInt. A small example is the test case of Trac #8963. Note [Product demands for function body] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ This example comes from shootout/binary_trees: Main.check' = \ b z ds. case z of z' { I# ip -> case ds_d13s of Main.Nil -> z' Main.Node s14k s14l s14m -> Main.check' (not b) (Main.check' b (case b { False -> I# (-# s14h s14k); True -> I# (+# s14h s14k) }) s14l) s14m } } } Here we *really* want to unbox z, even though it appears to be used boxed in the Nil case. Partly the Nil case is not a hot path. But more specifically, the whole function gets the CPR property if we do. So for the demand on the body of a RHS we use a product demand if it's a product type. %************************************************************************ %* * \subsection{Strictness signatures and types} %* * %************************************************************************ \begin{code}

unitVarDmd :: Var -> Demand -> DmdType unitVarDmd var dmd = DmdType (unitVarEnv var dmd) [] topRes addVarDmd :: DmdType -> Var -> Demand -> DmdType addVarDmd (DmdType fv ds res) var dmd = DmdType (extendVarEnv_C bothDmd fv var dmd) ds res addLazyFVs :: DmdType -> DmdEnv -> DmdType addLazyFVs dmd_ty lazy_fvs = dmd_ty `bothDmdType` mkBothDmdArg lazy_fvs -- Using bothDmdType (rather than just both'ing the envs) -- is vital. Consider -- let f = \x -> (x,y) -- in error (f 3) -- Here, y is treated as a lazy-fv of f, but we must `bothDmd` that L -- demand with the bottom coming up from 'error' -- -- I got a loop in the fixpointer without this, due to an interaction -- with the lazy_fv filtering in dmdAnalRhs. Roughly, it was -- letrec f n x -- = letrec g y = x `fatbar` -- letrec h z = z + ...g... -- in h (f (n-1) x) -- in ... -- In the initial iteration for f, f=Bot -- Suppose h is found to be strict in z, but the occurrence of g in its RHS -- is lazy. Now consider the fixpoint iteration for g, esp the demands it -- places on its free variables. Suppose it places none. Then the -- x `fatbar` ...call to h... -- will give a x->V demand for x. That turns into a L demand for x, -- which floats out of the defn for h. Without the modifyEnv, that -- L demand doesn't get both'd with the Bot coming up from the inner -- call to f. So we just get an L demand for x for g.\end{code} Note [do not strictify the argument dictionaries of a dfun] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The typechecker can tie recursive knots involving dfuns, so we do the conservative thing and refrain from strictifying a dfun's argument dictionaries. \begin{code}

annotateBndr :: AnalEnv -> DmdType -> Var -> (DmdType, Var) -- The returned env has the var deleted -- The returned var is annotated with demand info -- according to the result demand of the provided demand type -- No effect on the argument demands annotateBndr env dmd_ty var | isTyVar var = (dmd_ty, var) | otherwise = (dmd_ty', set_idDemandInfo env var dmd') where (dmd_ty', dmd) = peelFV dmd_ty var dmd' | gopt Opt_DictsStrict (ae_dflags env) -- We never want to strictify a recursive let. At the moment -- annotateBndr is only call for non-recursive lets; if that -- changes, we need a RecFlag parameter and another guard here. = strictifyDictDmd (idType var) dmd | otherwise = dmd annotateBndrs :: AnalEnv -> DmdType -> [Var] -> (DmdType, [Var]) annotateBndrs env = mapAccumR (annotateBndr env) annotateLamBndrs :: AnalEnv -> DFunFlag -> DmdType -> [Var] -> (DmdType, [Var]) annotateLamBndrs env args_of_dfun ty bndrs = mapAccumR annotate ty bndrs where annotate dmd_ty bndr | isId bndr = annotateLamIdBndr env args_of_dfun dmd_ty Many bndr | otherwise = (dmd_ty, bndr) annotateLamIdBndr :: AnalEnv -> DFunFlag -- is this lambda at the top of the RHS of a dfun? -> DmdType -- Demand type of body -> Count -- One-shot-ness of the lambda -> Id -- Lambda binder -> (DmdType, -- Demand type of lambda Id) -- and binder annotated with demand annotateLamIdBndr env arg_of_dfun dmd_ty one_shot id -- For lambdas we add the demand to the argument demands -- Only called for Ids = ASSERT( isId id ) -- pprTrace "annLamBndr" (vcat [ppr id, ppr _dmd_ty]) $ (final_ty, setOneShotness one_shot (set_idDemandInfo env id dmd')) where -- Watch out! See note [Lambda-bound unfoldings] final_ty = case maybeUnfoldingTemplate (idUnfolding id) of Nothing -> main_ty Just unf -> main_ty `bothDmdType` unf_ty where (unf_ty, _) = dmdAnalStar env dmd unf main_ty = addDemand dmd dmd_ty' (dmd_ty', dmd) = peelFV dmd_ty id dmd' | gopt Opt_DictsStrict (ae_dflags env), -- see Note [do not strictify the argument dictionaries of a dfun] not arg_of_dfun = strictifyDictDmd (idType id) dmd | otherwise = dmd deleteFVs :: DmdType -> [Var] -> DmdType deleteFVs (DmdType fvs dmds res) bndrs = DmdType (delVarEnvList fvs bndrs) dmds res\end{code} Note [CPR for sum types] ~~~~~~~~~~~~~~~~~~~~~~~~ At the moment we do not do CPR for let-bindings that * non-top level * bind a sum type Reason: I found that in some benchmarks we were losing let-no-escapes, which messed it all up. Example let j = \x. .... in case y of True -> j False False -> j True If we w/w this we get let j' = \x. .... in case y of True -> case j' False of { (# a #) -> Just a } False -> case j' True of { (# a #) -> Just a } Notice that j' is not a let-no-escape any more. However this means in turn that the *enclosing* function may be CPR'd (via the returned Justs). But in the case of sums, there may be Nothing alternatives; and that messes up the sum-type CPR. Conclusion: only do this for products. It's still not guaranteed OK for products, but sums definitely lose sometimes. Note [CPR for thunks] ~~~~~~~~~~~~~~~~~~~~~ If the rhs is a thunk, we usually forget the CPR info, because it is presumably shared (else it would have been inlined, and so we'd lose sharing if w/w'd it into a function). E.g. let r = case expensive of (a,b) -> (b,a) in ... If we marked r as having the CPR property, then we'd w/w into let $wr = \() -> case expensive of (a,b) -> (# b, a #) r = case $wr () of (# b,a #) -> (b,a) in ... But now r is a thunk, which won't be inlined, so we are no further ahead. But consider f x = let r = case expensive of (a,b) -> (b,a) in if foo r then r else (x,x) Does f have the CPR property? Well, no. However, if the strictness analyser has figured out (in a previous iteration) that it's strict, then we DON'T need to forget the CPR info. Instead we can retain the CPR info and do the thunk-splitting transform (see WorkWrap.splitThunk). This made a big difference to PrelBase.modInt, which had something like modInt = \ x -> let r = ... -> I# v in ...body strict in r... r's RHS isn't a value yet; but modInt returns r in various branches, so if r doesn't have the CPR property then neither does modInt Another case I found in practice (in Complex.magnitude), looks like this: let k = if ... then I# a else I# b in ... body strict in k .... (For this example, it doesn't matter whether k is returned as part of the overall result; but it does matter that k's RHS has the CPR property.) Left to itself, the simplifier will make a join point thus: let $j k = ...body strict in k... if ... then $j (I# a) else $j (I# b) With thunk-splitting, we get instead let $j x = let k = I#x in ...body strict in k... in if ... then $j a else $j b This is much better; there's a good chance the I# won't get allocated. The difficulty with this is that we need the strictness type to look at the body... but we now need the body to calculate the demand on the variable, so we can decide whether its strictness type should have a CPR in it or not. Simple solution: a) use strictness info from the previous iteration b) make sure we do at least 2 iterations, by doing a second round for top-level non-recs. Top level recs will get at least 2 iterations except for totally-bottom functions which aren't very interesting anyway. NB: strictly_demanded is never true of a top-level Id, or of a recursive Id. Note [Optimistic CPR in the "virgin" case] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Demand and strictness info are initialized by top elements. However, this prevents from inferring a CPR property in the first pass of the analyser, so we keep an explicit flag ae_virgin in the AnalEnv datatype. We can't start with 'not-demanded' (i.e., top) because then consider f x = let t = ... I# x in if ... then t else I# y else f x' In the first iteration we'd have no demand info for x, so assume not-demanded; then we'd get TopRes for f's CPR info. Next iteration we'd see that t was demanded, and so give it the CPR property, but by now f has TopRes, so it will stay TopRes. Instead, by checking the ae_virgin flag at the first time round, we say 'yes t is demanded' the first time. However, this does mean that for non-recursive bindings we must iterate twice to be sure of not getting over-optimistic CPR info, in the case where t turns out to be not-demanded. This is handled by dmdAnalTopBind. Note [NOINLINE and strictness] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The strictness analyser used to have a HACK which ensured that NOINLNE things were not strictness-analysed. The reason was unsafePerformIO. Left to itself, the strictness analyser would discover this strictness for unsafePerformIO: unsafePerformIO: C(U(AV)) But then consider this sub-expression unsafePerformIO (\s -> let r = f x in case writeIORef v r s of (# s1, _ #) -> (# s1, r #) The strictness analyser will now find that r is sure to be eval'd, and may then hoist it out. This makes tests/lib/should_run/memo002 deadlock. Solving this by making all NOINLINE things have no strictness info is overkill. In particular, it's overkill for runST, which is perfectly respectable. Consider f x = runST (return x) This should be strict in x. So the new plan is to define unsafePerformIO using the 'lazy' combinator: unsafePerformIO (IO m) = lazy (case m realWorld# of (# _, r #) -> r) Remember, 'lazy' is a wired-in identity-function Id, of type a->a, which is magically NON-STRICT, and is inlined after strictness analysis. So unsafePerformIO will look non-strict, and that's what we want. Now we don't need the hack in the strictness analyser. HOWEVER, this decision does mean that even a NOINLINE function is not entirely opaque: some aspect of its implementation leaks out, notably its strictness. For example, if you have a function implemented by an error stub, but which has RULES, you may want it not to be eliminated in favour of error! Note [Lazy and unleasheable free variables] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We put the strict and once-used FVs in the DmdType of the Id, so that at its call sites we unleash demands on its strict fvs. An example is 'roll' in imaginary/wheel-sieve2 Something like this: roll x = letrec go y = if ... then roll (x-1) else x+1 in go ms We want to see that roll is strict in x, which is because go is called. So we put the DmdEnv for x in go's DmdType. Another example: f :: Int -> Int -> Int f x y = let t = x+1 h z = if z==0 then t else if z==1 then x+1 else x + h (z-1) in h y Calling h does indeed evaluate x, but we can only see that if we unleash a demand on x at the call site for t. Incidentally, here's a place where lambda-lifting h would lose the cigar --- we couldn't see the joint strictness in t/x ON THE OTHER HAND We don't want to put *all* the fv's from the RHS into the DmdType, because that makes fixpointing very slow --- the DmdType gets full of lazy demands that are slow to converge. Note [Lamba-bound unfoldings] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We allow a lambda-bound variable to carry an unfolding, a facility that is used exclusively for join points; see Note [Case binders and join points]. If so, we must be careful to demand-analyse the RHS of the unfolding! Example \x. \y{=Just x}. Then if uses 'y', then transitively it uses 'x', and we must not forget that fact, otherwise we might make 'x' absent when it isn't. %************************************************************************ %* * \subsection{Strictness signatures} %* * %************************************************************************ \begin{code}

type DFunFlag = Bool -- indicates if the lambda being considered is in the -- sequence of lambdas at the top of the RHS of a dfun notArgOfDfun :: DFunFlag notArgOfDfun = False data AnalEnv = AE { ae_dflags :: DynFlags , ae_sigs :: SigEnv , ae_virgin :: Bool -- True on first iteration only -- See Note [Initialising strictness] , ae_rec_tc :: RecTcChecker , ae_fam_envs :: FamInstEnvs } -- We use the se_env to tell us whether to -- record info about a variable in the DmdEnv -- We do so if it's a LocalId, but not top-level -- -- The DmdEnv gives the demand on the free vars of the function -- when it is given enough args to satisfy the strictness signature type SigEnv = VarEnv (StrictSig, TopLevelFlag) instance Outputable AnalEnv where ppr (AE { ae_sigs = env, ae_virgin = virgin }) = ptext (sLit "AE") <+> braces (vcat [ ptext (sLit "ae_virgin =") <+> ppr virgin , ptext (sLit "ae_sigs =") <+> ppr env ]) emptyAnalEnv :: DynFlags -> FamInstEnvs -> AnalEnv emptyAnalEnv dflags fam_envs = AE { ae_dflags = dflags , ae_sigs = emptySigEnv , ae_virgin = True , ae_rec_tc = initRecTc , ae_fam_envs = fam_envs } emptySigEnv :: SigEnv emptySigEnv = emptyVarEnv sigEnv :: AnalEnv -> SigEnv sigEnv = ae_sigs updSigEnv :: AnalEnv -> SigEnv -> AnalEnv updSigEnv env sigs = env { ae_sigs = sigs } extendAnalEnv :: TopLevelFlag -> AnalEnv -> Id -> StrictSig -> AnalEnv extendAnalEnv top_lvl env var sig = env { ae_sigs = extendSigEnv top_lvl (ae_sigs env) var sig } extendSigEnv :: TopLevelFlag -> SigEnv -> Id -> StrictSig -> SigEnv extendSigEnv top_lvl sigs var sig = extendVarEnv sigs var (sig, top_lvl) lookupSigEnv :: AnalEnv -> Id -> Maybe (StrictSig, TopLevelFlag) lookupSigEnv env id = lookupVarEnv (ae_sigs env) id getStrictness :: AnalEnv -> Id -> StrictSig getStrictness env fn | isGlobalId fn = idStrictness fn | Just (sig, _) <- lookupSigEnv env fn = sig | otherwise = nopSig addInitialSigs :: TopLevelFlag -> AnalEnv -> [Id] -> AnalEnv -- See Note [Initialising strictness] addInitialSigs top_lvl env@(AE { ae_sigs = sigs, ae_virgin = virgin }) ids = env { ae_sigs = extendVarEnvList sigs [ (id, (init_sig id, top_lvl)) | id <- ids ] } where init_sig | virgin = \_ -> botSig | otherwise = idStrictness nonVirgin :: AnalEnv -> AnalEnv nonVirgin env = env { ae_virgin = False } extendSigsWithLam :: AnalEnv -> Id -> AnalEnv -- Extend the AnalEnv when we meet a lambda binder extendSigsWithLam env id | isId id , isStrictDmd (idDemandInfo id) || ae_virgin env -- See Note [Optimistic CPR in the "virgin" case] -- See Note [Initial CPR for strict binders] , Just (dc,_,_,_) <- deepSplitProductType_maybe (ae_fam_envs env) $ idType id = extendAnalEnv NotTopLevel env id (cprProdSig (dataConRepArity dc)) | otherwise = env set_idDemandInfo :: AnalEnv -> Id -> Demand -> Id set_idDemandInfo env id dmd = setIdDemandInfo id (zapDemand (ae_dflags env) dmd) set_idStrictness :: AnalEnv -> Id -> StrictSig -> Id set_idStrictness env id sig = setIdStrictness id (zapStrictSig (ae_dflags env) sig) dumpStrSig :: CoreProgram -> SDoc dumpStrSig binds = vcat (map printId ids) where ids = sortBy (stableNameCmp `on` getName) (concatMap getIds binds) getIds (NonRec i _) = [ i ] getIds (Rec bs) = map fst bs printId id | isExportedId id = ppr id <> colon <+> pprIfaceStrictSig (idStrictness id) | otherwise = empty\end{code} Note [Initial CPR for strict binders] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ CPR is initialized for a lambda binder in an optimistic manner, i.e, if the binder is used strictly and at least some of its components as a product are used, which is checked by the value of the absence demand. If the binder is marked demanded with a strict demand, then give it a CPR signature, because in the likely event that this is a lambda on a fn defn [we only use this when the lambda is being consumed with a call demand], it'll be w/w'd and so it will be CPR-ish. E.g. f = \x::(Int,Int). if ...strict in x... then x else (a,b) We want f to have the CPR property because x does, by the time f has been w/w'd Also note that we only want to do this for something that definitely has product type, else we may get over-optimistic CPR results (e.g. from \x -> x!). Note [Initialising strictness] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ See section 9.2 (Finding fixpoints) of the paper. Our basic plan is to initialise the strictness of each Id in a recursive group to "bottom", and find a fixpoint from there. However, this group B might be inside an *enclosing* recursiveb group A, in which case we'll do the entire fixpoint shebang on for each iteration of A. This can be illustrated by the following example: Example: f [] = [] f (x:xs) = let g [] = f xs g (y:ys) = y+1 : g ys in g (h x) At each iteration of the fixpoint for f, the analyser has to find a fixpoint for the enclosed function g. In the meantime, the demand values for g at each iteration for f are *greater* than those we encountered in the previous iteration for f. Therefore, we can begin the fixpoint for g not with the bottom value but rather with the result of the previous analysis. I.e., when beginning the fixpoint process for g, we can start from the demand signature computed for g previously and attached to the binding occurrence of g. To speed things up, we initialise each iteration of A (the enclosing one) from the result of the last one, which is neatly recorded in each binder. That way we make use of earlier iterations of the fixpoint algorithm. (Cunning plan.) But on the *first* iteration we want to *ignore* the current strictness of the Id, and start from "bottom". Nowadays the Id can have a current strictness, because interface files record strictness for nested bindings. To know when we are in the first iteration, we look at the ae_virgin field of the AnalEnv.