{- (c) The GRASP/AQUA Project, Glasgow University, 1993-1998 ----------------- A demand analysis ----------------- -} {-# LANGUAGE CPP #-} module DmdAnal ( dmdAnalProgram ) where #include "HsVersions.h" import GhcPrelude import DynFlags import WwLib ( findTypeShape, deepSplitProductType_maybe ) import Demand -- All of it import CoreSyn import CoreSeq ( seqBinds ) import Outputable import VarEnv import BasicTypes import Data.List import DataCon import Id import CoreUtils ( exprIsHNF, exprType, exprIsTrivial ) import TyCon import Type import Coercion ( Coercion, coVarsOfCo ) import FamInstEnv import Util import Maybes ( isJust ) import TysWiredIn import TysPrim ( realWorldStatePrimTy ) import ErrUtils ( dumpIfSet_dyn ) import Name ( getName, stableNameCmp ) import Data.Function ( on ) import UniqSet {- ************************************************************************ * * \subsection{Top level stuff} * * ************************************************************************ -} 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_str_signatures "Strictness signatures" $ dumpStrSig binds_plus_dmds ; -- See Note [Stamp out space leaks in demand analysis] seqBinds binds_plus_dmds `seq` 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 env (NonRec id rhs) = (extendAnalEnv TopLevel env id2 (idStrictness id2), NonRec id2 rhs2) where ( _, _, rhs1) = dmdAnalRhsLetDown TopLevel Nothing env cleanEvalDmd id rhs ( _, id2, rhs2) = dmdAnalRhsLetDown TopLevel Nothing (nonVirgin env) cleanEvalDmd id rhs1 -- Do two passes to improve CPR information -- See Note [CPR for thunks] -- See Note [Optimistic CPR in the "virgin" case] -- See Note [Initial CPR for strict binders] dmdAnalTopBind env (Rec pairs) = (env', Rec pairs') where (env', _, pairs') = dmdFix TopLevel env cleanEvalDmd pairs -- We get two iterations automatically -- c.f. the NonRec case above {- Note [Stamp out space leaks in demand analysis] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The demand analysis pass outputs a new copy of the Core program in which binders have been annotated with demand and strictness information. It's tiresome to ensure that this information is fully evaluated everywhere that we produce it, so we just run a single seqBinds over the output before returning it, to ensure that there are no references holding on to the input Core program. This makes a ~30% reduction in peak memory usage when compiling DynFlags (cf Trac #9675 and #13426). This is particularly important when we are doing late demand analysis, since we don't do a seqBinds at any point thereafter. Hence code generation would hold on to an extra copy of the Core program, via unforced thunks in demand or strictness information; and it is the most memory-intensive part of the compilation process, so this added seqBinds makes a big difference in peak memory usage. -} {- ************************************************************************ * * \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! -} -- 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 Cardinality Analysis paper dmdAnalStar :: AnalEnv -> Demand -- This one takes a *Demand* -> CoreExpr -> (BothDmdArg, CoreExpr) dmdAnalStar env dmd e | (defer_and_use, cd) <- toCleanDmd dmd (exprType e) , (dmd_ty, e') <- dmdAnal env cd e = (postProcessDmdType defer_and_use dmd_ty, e') -- Main Demand Analsysis machinery dmdAnal, 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 env d e = -- pprTrace "dmdAnal" (ppr d <+> ppr e) $ dmdAnal' env d e dmdAnal' _ _ (Lit lit) = (nopDmdType, Lit lit) dmdAnal' _ _ (Type ty) = (nopDmdType, Type ty) -- Doesn't happen, in fact dmdAnal' _ _ (Coercion co) = (unitDmdType (coercionDmdEnv co), Coercion co) dmdAnal' env dmd (Var var) = (dmdTransform env var dmd, Var var) dmdAnal' env dmd (Cast e co) = (dmd_ty `bothDmdType` mkBothDmdArg (coercionDmdEnv co), 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 -- Lots of the other code is there to make this -- beautiful, compositional, application rule :-) dmdAnal' env dmd (App fun arg) = -- This case handles value arguments (type args handled above) -- Crucially, coercions /are/ handled here, because they are -- value arguments (Trac #10288) let 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 @dmdAnalRhsLetDown@ 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) = peelCallDmd dmd -- body_dmd: a demand to analyze the body env' = extendSigsWithLam env var (body_ty, body') = dmdAnal env' body_dmd body (lam_ty, var') = annotateLamIdBndr env notArgOfDfun body_ty var in (postProcessUnsat defer_and_use lam_ty, Lam var' body') dmdAnal' env dmd (Case scrut case_bndr ty [(DataAlt dc, bndrs, rhs)]) -- Only one alternative with a product constructor | let tycon = dataConTyCon dc , isJust (isDataProductTyCon_maybe tycon) , Just rec_tc' <- checkRecTc (ae_rec_tc env) tycon = let env_w_tc = env { ae_rec_tc = rec_tc' } env_alt = extendEnvForProdAlt env_w_tc scrut case_bndr dc bndrs (rhs_ty, rhs') = dmdAnal env_alt dmd rhs (alt_ty1, dmds) = findBndrsDmds env rhs_ty bndrs (alt_ty2, case_bndr_dmd) = findBndrDmd env False alt_ty1 case_bndr id_dmds = addCaseBndrDmd case_bndr_dmd dmds alt_ty3 | io_hack_reqd scrut dc bndrs = deferAfterIO alt_ty2 | otherwise = alt_ty2 -- Compute demand on the scrutinee -- See Note [Demand on scrutinee of a product case] scrut_dmd = mkProdDmd (addDataConStrictness dc id_dmds) (scrut_ty, scrut') = dmdAnal env scrut_dmd scrut res_ty = alt_ty3 `bothDmdType` toBothDmdArg scrut_ty case_bndr' = setIdDemandInfo case_bndr case_bndr_dmd bndrs' = setBndrsDemandInfo bndrs id_dmds 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_ty2 -- , text "res_ty" <+> ppr res_ty ]) $ (res_ty, Case scrut' case_bndr' ty [(DataAlt dc, bndrs', rhs')]) dmdAnal' env dmd (Case scrut case_bndr ty alts) = let -- Case expression with multiple alternatives (alt_tys, alts') = mapAndUnzip (dmdAnalAlt env dmd case_bndr) alts (scrut_ty, scrut') = dmdAnal env cleanEvalDmd scrut (alt_ty, case_bndr') = annotateBndr env (foldr lubDmdType botDmdType alt_tys) case_bndr -- NB: Base case is botDmdType, for empty case alternatives -- This is a unit for lubDmdType, and the right result -- when there really are no alternatives 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') -- Let bindings can be processed in two ways: -- Down (RHS before body) or Up (body before RHS). -- The following case handle the up variant. -- -- It is very simple. For let x = rhs in body -- * Demand-analyse 'body' in the current environment -- * Find the demand, 'rhs_dmd' placed on 'x' by 'body' -- * Demand-analyse 'rhs' in 'rhs_dmd' -- -- This is used for a non-recursive local let without manifest lambdas. -- This is the LetUp rule in the paper “Higher-Order Cardinality Analysis”. dmdAnal' env dmd (Let (NonRec id rhs) body) | useLetUp id rhs , Nothing <- unpackTrivial rhs -- dmdAnalRhsLetDown treats trivial right hand sides specially -- so if we have a trival right hand side, fall through to that. = (final_ty, Let (NonRec id' rhs') body') where (body_ty, body') = dmdAnal env dmd body (body_ty', id_dmd) = findBndrDmd env notArgOfDfun body_ty id id' = setIdDemandInfo id id_dmd (rhs_ty, rhs') = dmdAnalStar env (dmdTransformThunkDmd rhs id_dmd) rhs final_ty = body_ty' `bothDmdType` rhs_ty dmdAnal' env dmd (Let (NonRec id rhs) body) = (body_ty2, Let (NonRec id2 rhs') body') where (lazy_fv, id1, rhs') = dmdAnalRhsLetDown NotTopLevel Nothing env dmd id rhs env1 = extendAnalEnv NotTopLevel env id1 (idStrictness id1) (body_ty, body') = dmdAnal env1 dmd body (body_ty1, id2) = annotateBndr env body_ty id1 body_ty2 = addLazyFVs body_ty1 lazy_fv -- see Note [Lazy and unleashable free variables] -- 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 dmd pairs (body_ty, body') = dmdAnal env' dmd body body_ty1 = deleteFVs body_ty (map fst pairs) body_ty2 = addLazyFVs body_ty1 lazy_fv -- see Note [Lazy and unleashable free variables] in body_ty2 `seq` (body_ty2, Let (Rec pairs') body') io_hack_reqd :: CoreExpr -> DataCon -> [Var] -> Bool -- See Note [IO hack in the demand analyser] io_hack_reqd scrut con bndrs | (bndr:_) <- bndrs , con == tupleDataCon Unboxed 2 , idType bndr `eqType` realWorldStatePrimTy , (fun, _) <- collectArgs scrut = case fun of Var f -> not (isPrimOpId f) _ -> True | otherwise = False dmdAnalAlt :: AnalEnv -> CleanDemand -> Id -> Alt Var -> (DmdType, Alt Var) dmdAnalAlt env dmd case_bndr (con,bndrs,rhs) | null bndrs -- Literals, DEFAULT, and nullary constructors , (rhs_ty, rhs') <- dmdAnal env dmd rhs = (rhs_ty, (con, [], rhs')) | otherwise -- Non-nullary data constructors , (rhs_ty, rhs') <- dmdAnal env dmd rhs , (alt_ty, dmds) <- findBndrsDmds env rhs_ty bndrs , let case_bndr_dmd = findIdDemand alt_ty case_bndr id_dmds = addCaseBndrDmd case_bndr_dmd dmds = (alt_ty, (con, setBndrsDemandInfo bndrs id_dmds, rhs')) {- 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'? Often not! If foo x s performs some observable action (including raising an exception with raiseIO#, modifying a mutable variable, or even ending the program normally), then we must not force 'y' (which may fail to terminate) until we have performed foo x s. 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 } However, consider f x s = case getMaskingState# s of (# s, r #) -> case x of I# x2 -> ... Here it is terribly sad to make 'f' lazy in 's'. After all, getMaskingState# is not going to diverge or throw an exception! This situation actually arises in GHC.IO.Handle.Internals.wantReadableHandle (on an MVar not an Int), and made a material difference. So if the scrutinee is a primop call, we *don't* apply the state hack: - If it is a simple, terminating one like getMaskingState, applying the hack is over-conservative. - If the primop is raise# then it returns bottom, so the case alternatives are already discarded. - If the primop can raise a non-IO exception, like divide by zero or seg-fault (eg writing an array out of bounds) then we don't mind evaluating 'x' first. Note [Demand on the scrutinee of a product case] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ When figuring out the demand on the scrutinee of a product case, we use the demands of the case alternative, i.e. id_dmds. But note that these include the demand on the case binder; see Note [Demand on case-alternative binders] in Demand.hs. This is crucial. Example: 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. 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|. ************************************************************************ * * Demand transformer * * ************************************************************************ -} 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 = unitDmdType (unitVarEnv var (mkOnceUsedDmd dmd)) {- ************************************************************************ * * \subsection{Bindings} * * ************************************************************************ -} -- Recursive bindings dmdFix :: TopLevelFlag -> AnalEnv -- Does not include bindings for this binding -> CleanDemand -> [(Id,CoreExpr)] -> (AnalEnv, DmdEnv, [(Id,CoreExpr)]) -- Binders annotated with stricness info dmdFix top_lvl env let_dmd orig_pairs = loop 1 initial_pairs where bndrs = map fst orig_pairs -- See Note [Initialising strictness] initial_pairs | ae_virgin env = [(setIdStrictness id botSig, rhs) | (id, rhs) <- orig_pairs ] | otherwise = orig_pairs -- If fixed-point iteration does not yield a result we use this instead -- See Note [Safe abortion in the fixed-point iteration] abort :: (AnalEnv, DmdEnv, [(Id,CoreExpr)]) abort = (env, lazy_fv', zapped_pairs) where (lazy_fv, pairs') = step True (zapIdStrictness orig_pairs) -- Note [Lazy and unleashable free variables] non_lazy_fvs = plusVarEnvList $ map (strictSigDmdEnv . idStrictness . fst) pairs' lazy_fv' = lazy_fv `plusVarEnv` mapVarEnv (const topDmd) non_lazy_fvs zapped_pairs = zapIdStrictness pairs' -- The fixed-point varies the idStrictness field of the binders, and terminates if that -- annotation does not change any more. loop :: Int -> [(Id,CoreExpr)] -> (AnalEnv, DmdEnv, [(Id,CoreExpr)]) loop n pairs | found_fixpoint = (final_anal_env, lazy_fv, pairs') | n == 10 = abort | otherwise = loop (n+1) pairs' where found_fixpoint = map (idStrictness . fst) pairs' == map (idStrictness . fst) pairs first_round = n == 1 (lazy_fv, pairs') = step first_round pairs final_anal_env = extendAnalEnvs top_lvl env (map fst pairs') step :: Bool -> [(Id, CoreExpr)] -> (DmdEnv, [(Id, CoreExpr)]) step first_round pairs = (lazy_fv, pairs') where -- In all but the first iteration, delete the virgin flag start_env | first_round = env | otherwise = nonVirgin env start = (extendAnalEnvs top_lvl start_env (map fst pairs), emptyDmdEnv) ((_,lazy_fv), pairs') = mapAccumL my_downRhs start 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 (lazy_fv1, id', rhs') = dmdAnalRhsLetDown top_lvl (Just bndrs) env let_dmd id rhs lazy_fv' = plusVarEnv_C bothDmd lazy_fv lazy_fv1 env' = extendAnalEnv top_lvl env id (idStrictness id') zapIdStrictness :: [(Id, CoreExpr)] -> [(Id, CoreExpr)] zapIdStrictness pairs = [(setIdStrictness id nopSig, rhs) | (id, rhs) <- pairs ] {- Note [Safe abortion in the fixed-point iteration] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Fixed-point iteration may fail to terminate. But we cannot simply give up and return the environment and code unchanged! We still need to do one additional round, for two reasons: * To get information on used free variables (both lazy and strict!) (see Note [Lazy and unleashable free variables]) * To ensure that all expressions have been traversed at least once, and any left-over strictness annotations have been updated. This final iteration does not add the variables to the strictness signature environment, which effectively assigns them 'nopSig' (see "getStrictness") -} -- Trivial RHS -- See Note [Demand analysis for trivial right-hand sides] dmdAnalTrivialRhs :: AnalEnv -> Id -> CoreExpr -> Var -> (DmdEnv, Id, CoreExpr) dmdAnalTrivialRhs env id rhs fn = (fn_fv, set_idStrictness env id fn_str, rhs) where 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 -- Let bindings can be processed in two ways: -- Down (RHS before body) or Up (body before RHS). -- dmdAnalRhsLetDown implements the Down variant: -- * assuming a demand of <L,U> -- * looking at the definition -- * determining a strictness signature -- -- It is used for toplevel definition, recursive definitions and local -- non-recursive definitions that have manifest lambdas. -- Local non-recursive definitions without a lambda are handled with LetUp. -- -- This is the LetDown rule in the paper “Higher-Order Cardinality Analysis”. dmdAnalRhsLetDown :: TopLevelFlag -> Maybe [Id] -- Just bs <=> recursive, Nothing <=> non-recursive -> AnalEnv -> CleanDemand -> Id -> CoreExpr -> (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. dmdAnalRhsLetDown top_lvl rec_flag env let_dmd id rhs | Just fn <- unpackTrivial rhs -- See Note [Demand analysis for trivial right-hand sides] = dmdAnalTrivialRhs env id rhs fn | otherwise = (lazy_fv, id', mkLams bndrs' body') where (bndrs, body, body_dmd) = case isJoinId_maybe id of Just join_arity -- See Note [Demand analysis for join points] | (bndrs, body) <- collectNBinders join_arity rhs -> (bndrs, body, let_dmd) Nothing | (bndrs, body) <- collectBinders rhs -> (bndrs, body, mkBodyDmd env body) 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 [Aggregated demand for cardinality] rhs_fv1 = case rec_flag of Just bs -> reuseEnv (delVarEnvList rhs_fv bs) Nothing -> rhs_fv -- See Note [Lazy and unleashable free variables] (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 (isJoinId id) 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] mkBodyDmd :: AnalEnv -> CoreExpr -> CleanDemand -- See Note [Product demands for function body] mkBodyDmd env body = case deepSplitProductType_maybe (ae_fam_envs env) (exprType body) of Nothing -> cleanEvalDmd Just (dc, _, _, _) -> cleanEvalProdDmd (dataConRepArity dc) 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 -- | If given the RHS of a let-binding, this 'useLetUp' determines -- whether we should process the binding up (body before rhs) or -- down (rhs before body). -- -- We use LetDown if there is a chance to get a useful strictness signature. -- This is the case when there are manifest value lambdas or the binding is a -- join point (hence always acts like a function, not a value). useLetUp :: Var -> CoreExpr -> Bool useLetUp f _ | isJoinId f = False useLetUp f (Lam v e) | isTyVar v = useLetUp f e useLetUp _ (Lam _ _) = False useLetUp _ _ = True {- Note [Demand analysis for join points] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider g :: (Int,Int) -> Int g (p,q) = p+q f :: T -> Int -> Int f x p = g (join j y = (p,y) in case x of A -> j 3 B -> j 4 C -> (p,7)) If j was a vanilla function definition, we'd analyse its body with evalDmd, and think that it was lazy in p. But for join points we can do better! We know that j's body will (if called at all) be evaluated with the demand that consumes the entire join-binding, in this case the argument demand from g. Whizzo! g evaluates both components of its argument pair, so p will certainly be evaluated if j is called. For f to be strict in p, we need /all/ paths to evaluate p; in this case the C branch does so too, so we are fine. So, as usual, we need to transport demands on free variables to the call site(s). Compare Note [Lazy and unleashable free variables]. The implementation is easy. When analysing a join point, we can analyse its body with the demand from the entire join-binding (written let_dmd here). Another win for join points! Trac #13543. 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} * * ************************************************************************ -} unitDmdType :: DmdEnv -> DmdType unitDmdType dmd_env = DmdType dmd_env [] topRes coercionDmdEnv :: Coercion -> DmdEnv coercionDmdEnv co = mapVarEnv (const topDmd) (getUniqSet $ coVarsOfCo co) -- The VarSet from coVarsOfCo is really a VarEnv Var 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 dmdAnalRhsLetDown. 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. {- 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. -} setBndrsDemandInfo :: [Var] -> [Demand] -> [Var] setBndrsDemandInfo (b:bs) (d:ds) | isTyVar b = b : setBndrsDemandInfo bs (d:ds) | otherwise = setIdDemandInfo b d : setBndrsDemandInfo bs ds setBndrsDemandInfo [] ds = ASSERT( null ds ) [] setBndrsDemandInfo bs _ = pprPanic "setBndrsDemandInfo" (ppr bs) 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 | isId var = (dmd_ty', setIdDemandInfo var dmd) | otherwise = (dmd_ty, var) where (dmd_ty', dmd) = findBndrDmd env False dmd_ty var 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 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 -> Id -- Lambda binder -> (DmdType, -- Demand type of lambda Id) -- and binder annotated with demand annotateLamIdBndr env arg_of_dfun dmd_ty 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, setIdDemandInfo 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) = findBndrDmd env arg_of_dfun dmd_ty id deleteFVs :: DmdType -> [Var] -> DmdType deleteFVs (DmdType fvs dmds res) bndrs = DmdType (delVarEnvList fvs bndrs) dmds res {- 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 unleashable 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 * it makes the strictness signatures larger, and hence slows down fixpointing and * it is useless information at the call site anyways: For lazy, used-many times fv's we will never get any better result than that, no matter how good the actual demand on the function at the call site is (unless it is always absent, but then the whole binder is useless). Therefore we exclude lazy multiple-used fv's from the environment in the DmdType. But now the signature lies! (Missing variables are assumed to be absent.) To make up for this, the code that analyses the binding keeps the demand on those variable separate (usually called "lazy_fv") and adds it to the demand of the whole binding later. What if we decide _not_ to store a strictness signature for a binding at all, as we do when aborting a fixed-point iteration? The we risk losing the information that the strict variables are being used. In that case, we take all free variables mentioned in the (unsound) strictness signature, conservatively approximate the demand put on them (topDmd), and add that to the "lazy_fv" returned by "dmdFix". Note [Lambda-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}. <body> Then if <body> 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} * * ************************************************************************ -} 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 }) = text "AE" <+> braces (vcat [ text "ae_virgin =" <+> ppr virgin , text "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 -- | Extend an environment with the strictness IDs attached to the id extendAnalEnvs :: TopLevelFlag -> AnalEnv -> [Id] -> AnalEnv extendAnalEnvs top_lvl env vars = env { ae_sigs = extendSigEnvs top_lvl (ae_sigs env) vars } extendSigEnvs :: TopLevelFlag -> SigEnv -> [Id] -> SigEnv extendSigEnvs top_lvl sigs vars = extendVarEnvList sigs [ (var, (idStrictness var, top_lvl)) | var <- vars] 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 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 extendEnvForProdAlt :: AnalEnv -> CoreExpr -> Id -> DataCon -> [Var] -> AnalEnv -- See Note [CPR in a product case alternative] extendEnvForProdAlt env scrut case_bndr dc bndrs = foldl do_con_arg env1 ids_w_strs where env1 = extendAnalEnv NotTopLevel env case_bndr case_bndr_sig ids_w_strs = filter isId bndrs `zip` dataConRepStrictness dc case_bndr_sig = cprProdSig (dataConRepArity dc) fam_envs = ae_fam_envs env do_con_arg env (id, str) | let is_strict = isStrictDmd (idDemandInfo id) || isMarkedStrict str , ae_virgin env || (is_var_scrut && is_strict) -- See Note [CPR in a product case alternative] , Just (dc,_,_,_) <- deepSplitProductType_maybe fam_envs $ idType id = extendAnalEnv NotTopLevel env id (cprProdSig (dataConRepArity dc)) | otherwise = env is_var_scrut = is_var scrut is_var (Cast e _) = is_var e is_var (Var v) = isLocalId v is_var _ = False addDataConStrictness :: DataCon -> [Demand] -> [Demand] -- See Note [Add demands for strict constructors] addDataConStrictness con ds = ASSERT2( equalLength strs ds, ppr con $$ ppr strs $$ ppr ds ) zipWith add ds strs where strs = dataConRepStrictness con add dmd str | isMarkedStrict str , not (isAbsDmd dmd) = dmd `bothDmd` seqDmd | otherwise = dmd findBndrsDmds :: AnalEnv -> DmdType -> [Var] -> (DmdType, [Demand]) -- Return the demands on the Ids in the [Var] findBndrsDmds env dmd_ty bndrs = go dmd_ty bndrs where go dmd_ty [] = (dmd_ty, []) go dmd_ty (b:bs) | isId b = let (dmd_ty1, dmds) = go dmd_ty bs (dmd_ty2, dmd) = findBndrDmd env False dmd_ty1 b in (dmd_ty2, dmd : dmds) | otherwise = go dmd_ty bs findBndrDmd :: AnalEnv -> Bool -> DmdType -> Id -> (DmdType, Demand) -- See Note [Trimming a demand to a type] in Demand.hs findBndrDmd env arg_of_dfun dmd_ty id = (dmd_ty', dmd') where dmd' = killUsageDemand (ae_dflags env) $ strictify $ trimToType starting_dmd (findTypeShape fam_envs id_ty) (dmd_ty', starting_dmd) = peelFV dmd_ty id id_ty = idType id strictify 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. , not arg_of_dfun -- See Note [Do not strictify the argument dictionaries of a dfun] = strictifyDictDmd id_ty dmd | otherwise = dmd fam_envs = ae_fam_envs env set_idStrictness :: AnalEnv -> Id -> StrictSig -> Id set_idStrictness env id sig = setIdStrictness id (killUsageSig (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 {- Note [CPR in a product case alternative] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ In a case alternative for a product type, we want to give some of the binders the CPR property. Specifically * The case binder; inside the alternative, the case binder always 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 By giving 'y' the CPR property, we ensure that 'f' does too, so we get 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 Of course there is the usual risk of re-boxing: we have 'x' available boxed and unboxed, but we return the unboxed version for the wrapper to box. If the wrapper doesn't cancel with its caller, we'll end up re-boxing something that we did have available in boxed form. * Any strict binders with product type, can use Note [Initial CPR for strict binders]. But we can go a little further. Consider data T = MkT !Int Int f2 (MkT x y) | y>0 = f2 (MkT x (y-1)) | otherwise = x For $wf2 we are going to unbox the MkT *and*, since it is strict, the first argument of the MkT; see Note [Add demands for strict constructors]. But then we don't want box it up again when returning it! We want 'f2' to have the CPR property, so we give 'x' the CPR property. * It's a bit delicate because if this case is scrutinising something other than an argument the original function, we really don't have the unboxed version available. E.g g v = case foo v of MkT x y | y>0 -> ... | otherwise -> x Here we don't have the unboxed 'x' available. Hence the is_var_scrut test when making use of the strictness annotation. Slightly ad-hoc, because even if the scrutinee *is* a variable it might not be a onre of the arguments to the original function, or a sub-component thereof. But it's simple, and nothing terrible happens if we get it wrong. e.g. Trac #10694. 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 'go' 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! We achieve the effect using addDataConStrictness. It is called at a case expression, such as the pattern match on (X a) in the example above. After computing how 'a' is used in the alternatives, we add an extra 'seqDmd' to it. The case alternative isn't itself strict in the sub-components, but simply evaluating the scrutinee to HNF does force those sub-components. If the argument is not used at all in the alternative (i.e. it is Absent), then *don't* add a 'seqDmd'. If we do, it makes it look used and hence it'll be passed to the worker when it doesn't need to be. Hence the isAbsDmd test in addDataConStrictness. 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! 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. Here's a concrete example ('f1' in test T10482a), assuming h is strict: f1 :: Int -> Int f1 x = case h x of A -> x B -> f1 (x-1) C -> x+1 If we notice that 'x' is used strictly, we can give it the CPR property; and hence f1 gets the CPR property too. It's sound (doesn't change strictness) to give it the CPR property because by the time 'x' is returned (case A above), it'll have been evaluated (by the wrapper of 'h' in the example). Moreover, if f itself is strict in x, then we'll pass x unboxed to f1, and so the boxed version *won't* be available; in that case it's very helpful to give 'x' the CPR property. 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!). * See Note [CPR examples] Note [CPR examples] ~~~~~~~~~~~~~~~~~~~~ Here are some examples (stranal/should_compile/T10482a) of the usefulness of Note [CPR in a product case alternative]. The main point: all of these functions can have the CPR property. ------- f1 ----------- -- x is used strictly by h, so it'll be available -- unboxed before it is returned in the True branch f1 :: Int -> Int f1 x = case h x x of True -> x False -> f1 (x-1) ------- f2 ----------- -- x is a strict field of MkT2, so we'll pass it unboxed -- to $wf2, so it's available unboxed. This depends on -- the case expression analysing (a subcomponent of) one -- of the original arguments to the function, so it's -- a bit more delicate. data T2 = MkT2 !Int Int f2 :: T2 -> Int f2 (MkT2 x y) | y>0 = f2 (MkT2 x (y-1)) | otherwise = x ------- f3 ----------- -- h is strict in x, so x will be unboxed before it -- is rerturned in the otherwise case. data T3 = MkT3 Int Int f1 :: T3 -> Int f1 (MkT3 x y) | h x y = f3 (MkT3 x (y-1)) | otherwise = x ------- f4 ----------- -- Just like f2, but MkT4 can't unbox its strict -- argument automatically, as f2 can data family Foo a newtype instance Foo Int = Foo Int data T4 a = MkT4 !(Foo a) Int f4 :: T4 Int -> Int f4 (MkT4 x@(Foo v) y) | y>0 = f4 (MkT4 x (y-1)) | otherwise = v 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* recursive 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. Note [Final Demand Analyser run] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Some of the information that the demand analyser determines is not always preserved by the simplifier. For example, the simplifier will happily rewrite \y [Demand=1*U] let x = y in x + x to \y [Demand=1*U] y + y which is quite a lie. The once-used information is (currently) only used by the code generator, though. So: * We zap the used-once info in the worker-wrapper; see Note [Zapping Used Once info in WorkWrap] in WorkWrap. If it's not reliable, it's better not to have it at all. * Just before TidyCore, we add a pass of the demand analyser, but WITHOUT subsequent worker/wrapper and simplifier, right before TidyCore. See SimplCore.getCoreToDo. This way, correct information finds its way into the module interface (strictness signatures!) and the code generator (single-entry thunks!) Note that, in contrast, the single-call information (C1(..)) /can/ be relied upon, as the simplifier tends to be very careful about not duplicating actual function calls. Also see #11731. -}