## Abstract

In his classic monograph, *Social Choice and Individual Values*, Arrow introduced the notion of a decisive coalition of voters as part of his mathematical framework for social choice theory. The subsequent literature on Arrow’s Impossibility Theorem has shown the importance for social choice theory of reasoning about coalitions of voters with different grades of decisiveness. The goal of this paper is a fine-grained analysis of reasoning about decisive coalitions, formalizing how the concept of a decisive coalition gives rise to a social choice theoretic language and logic all of its own. We show that given Arrow’s axioms of the Independence of Irrelevant Alternatives and Universal Domain, rationality postulates for social preference correspond to strong axioms about decisive coalitions. We demonstrate this correspondence with results of a kind familiar in economics—representation theorems—as well as results of a kind coming from mathematical logic—completeness theorems. We present a complete logic for reasoning about decisive coalitions, along with formal proofs of Arrow’s and Wilson’s theorems. In addition, we prove the correctness of an algorithm for calculating, given any social rationality postulate of a certain form in the language of binary preference, the corresponding axiom in the language of decisive coalitions. These results suggest for social choice theory new perspectives and tools from logic.

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## Notes

For instance, see Saari (2008) for a geometric perspective, Baigent (2010) for a topological perspective, Abramsky (2015) for a category theoretic perspective, Lauwers and Van Liedekerke (1995) and Herzberg and Eckert (2012) for model-theoretic perspectives, Fishburn and Rubinstein (1986) and Shelah (2005) for algebraic perspectives, and Bao and Halpern (2017) for a quantum perspective.

Reasoning about decisive coalitions is also central in the theory of judgment aggregation (see, e.g., List and Polak 2010).

In a chapter entitled “Origins of the Impossibility Theorem” (Arrow 2014), Arrow recalls: “As it happens, during my college years, I was fascinated by mathematical logic, a subject I read on my own until, by a curious set of chances, the great Polish logician, Alfred Tarski, taught one year at The City College (in New York), where I was a senior. He chose to give a course on the calculus of relations, and I was introduced to such topics as transitivity and orderings” (p. 154). Sen (2017, p. xvi) even reports having taught mathematical logic at Delhi University. For his views on the relation between logic and economics, see Sen (2017, p. 108).

Earlier Murakami (1968) applied results about three-valued logic to the analysis of voting rules.

This separation of the social choice theoretic content and the combinatorial content of Arrow’s theorem is also a theme in Makinson (1996).

Strictly speaking, as Sen—like Arrow—takes the weak preference relation

*R*as primitive, Sen’s CCRs send a vector of reflexive, transitive, and complete relations on*X*to a binary relation on*X*.SPA is not as standard as the other conditions, but it has a natural connection to our main topic of decisive coalitions, explained at the end of this section.

For the latter implication, if \(\mathbf {P},\mathbf {P}'\in L(X)\), then \(\mathbf {P}(x^\star ,y)=\mathbf {P}'(x^\star ,y)\) implies \(\mathbf {P}(y, x^\star )=\mathbf {P}'(y, x^\star )\) and hence \(\mathbf {P}_{\mid \{x^\star ,y\}}=\mathbf {P}'_{\mid \{x^\star ,y\}}\). So if \(\mathrm {dom}(f)=L(X)\), then the assumption of SPA implies that for all \(x,y\in X\setminus \{x^\star \}\), \(\mathbf {P}_{\mid \{x,y\}}=\mathbf {P}'_{\mid \{x,y\}}\) and \(\mathbf {P}_{\mid \{x^\star ,y\}}=\mathbf {P}'_{\mid \{x^\star ,y\}}\), which implies \(\mathbf {P}=\mathbf {P}'\), so \(f(\mathbf {P})=f(\mathbf {P}')\).

In Taylor and Zwicker (1999), the definition of a ‘simple game’ requires that

*W*be closed under supersets, while the term ‘hypergraph’ is used for the more general structures.The reason for this finiteness assumption is given in footnote 18.

Note that ‘\(\equiv \)’ is a symbol in our formal language, which will be interpreted using the equality relation (Definition 4.4), which we denote by ‘\(=\)’.

We use ‘\({:}{=}\)’ to mean “is by definition an abbreviation for.”

A different approach would allow distinct alternative labels to stand for the same alternative, but would then add identity formulas \(x= y\) to our language, so that we could express a requirement of distinctness when needed by a formula of the form \(\lnot \;x =y\).

In fact, we also have all the resources to show that the answer is

*yes*for the class of CCRs satisfying LD, IIA, and TR (resp. FR), but we do not want to overload the reader with variations on our results.The axiom \(\lnot (0\equiv 1)\) reflects our assumption that the set

*V*of all voters is nonempty.A lower bound is easy: since our logic includes propositional logic, deciding theoremhood is a co-NP-hard problem. The question is: what is an upper bound?

Note that here we use the assumption that \(\mathsf {Alt}\) is finite (which follows from our assumptions that

*X*is finite and that \(|\mathsf {Alt}|=|X|\)), so that*EA*can be written as a finite conjunction in our language.For \(\widehat{\mathbf {T}}\), we can drop the assumption that \(x\ne v\) for part 1 and the assumption that \(w\ne y\) for part 2; and we can drop the same assumptions for \(\overline{\mathbf {T}}\) if we add the assumption \(\lnot (a\equiv 0)\) to the conditionals.

By ‘Boolean reasoning’, we mean a combined use of valid Boolean equations, Leibniz’s law, and propositional logic.

*A*is*semi-decisive for**x**over**y**according to**f*if and only if for any \(\mathbf {P}\in \mathrm {dom}(f)\), if \(A\subseteq \mathbf {P}(x,y)\), then*not*\(y f(\mathbf {P})x\).By \(WEA(\mathbf 1 )\), we mean the

*WEA*formula of Sect. 5.1 in which each \(c_{x,y}\) is 1.To make the following argument fully rigorous, we need to define the formal syntax of the language in which these postulates are given, as well as the formal syntax of the language into which we are translating them below. The former can be the language of predicate logic with a single binary predicate symbols

*P*, and the latter can be the language of (almost) decisiveness in Sect. 4.Thanks to Mikayla Kelley for this way of distinguishing the logic of decisive coalitions for Condorcet and Borda.

This is because there can be CCRs

*f*and \(f'\) such that:*f*and \(f'\) have the same (almost) decisive coalitions, i.e., for all \(x,y\in X\), \(\widehat{D}_f(x,y)=\widehat{D}_{f'}(x,y)\) and \(\overline{D}_f(x,y)=\overline{D}_{f'}(x,y)\);*f*does not satisfy non-null; and \(f'\) does satisfy non-null. For example, let*f*be the CCR such that for all profiles \(\mathbf {P}\) and \(x,y\in X\),*not*\(xf(\mathbf {P})y\). Given \(x,y\in X\), we will define a CCR \(f_{x,y}\) that differs from*f*only on the profile \(\mathbf {P}^\star \) such that for all \(x,y\in X\) and \(i\in V\),*not*\(x P_i y\); for this \(\mathbf {P}^\star \), let \(z f_{x,y}(\mathbf {P}^\star ) w\) if and only if either \(z=x\) and \(w\ne x\), or \(z\ne y\) and \(w=y\), so \(f_{x,y}(\mathbf {P}^\star )\) is a strict weak order. Then for all \(z,w\in X\), if \(z\ne w\), then \(\widehat{D}_f(z,w)=\widehat{D}_{f_{x,y}}(z,w)= \overline{D}_f(z,w)=\overline{D}_{f_{x,y}}(z,w)=\varnothing \), while if \(z=w\), then \(\widehat{D}_f(z,w)=\widehat{D}_{f_{x,y}}(z,w)=\wp (V)\) and \(\overline{D}_f(z,w)=\overline{D}_{f_{x,y}}(z,w)=\wp (V)\setminus \{\varnothing \}\);*f*does not satisfy non-null; and \(f_{x,y}\) does satisfy non-null.

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## Author information

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### Corresponding author

## Additional information

We wish to thank the referees for *Social Choice and Welfare*, Mikayla Kelley, and Kotaro Suzumura for helpful comments.

## Appendix

### Appendix

In this Appendix, we sketch the proofs of the completeness theorems stated in Theorems 4.8.2, 4.8.4, and 7.4.2. A similar approach can be used to prove completeness theorems for logics that include inverse decisiveness as in Sect. 7.4.

For a logic \(\mathbf {L}\), a set \(\Sigma \) of formulas is \(\mathbf {L}\)-*inconsistent* if and only if there are \(\sigma _1,\dots ,\sigma _n\in \Sigma \) such that \(\vdash _\mathbf {L} (\sigma _1\wedge \dots \wedge \sigma _n)\rightarrow \bot \), where \(\bot \) is any formula of the form \(\varphi \wedge \lnot \varphi \). Otherwise \(\Sigma \) is \(\mathbf {L}\)-*consistent*. We say that a formula \(\varphi \) is \(\mathbf {L}\)-consistent if and only if \(\{\varphi \}\) is \(\mathbf {L}\)-consistent. Finally, a set \(\Sigma \) of formulas is *maximally*\(\mathbf {L}\)-*consistent* if and only if for every set \(\Delta \) of formulas, \(\Sigma \subsetneq \Delta \) implies that \(\Delta \) is \(\mathbf {L}\)-inconsistent.

To say that a logic \(\mathbf {L}\) is complete with respect to a class *K* of CCRs is to say that for any formula \(\psi \), if \(\psi \) is true of all CCRs from *K*, then \(\psi \) is a theorem of \(\mathbf {L}\). Equivalently, if \(\psi \) is *not* a theorem of \(\mathbf {L}\), then there is a CCR from *K* of which \(\psi \) is *not* true. Since by propositional logic, we have \(\not \vdash _\mathbf {L} \psi \) if and only if \(\not \vdash _\mathbf {L} \lnot \psi \rightarrow \bot \), the formula \(\psi \) not being a theorem of \(\mathbf {L}\) is equivalent to \(\lnot \psi \) being \(\mathbf {L}\)-consistent. Moreover, there being a CCR from *K* of which \(\psi \) is not true is equivalent to there being a CCR from *K* of which \(\lnot \psi \) is true. Thus, to prove that \(\mathbf {L}\) is complete with respect to *K*, it suffices to show that for any formula \(\varphi \), if \(\varphi \) is \(\mathbf {L}\)-consistent, then there is a CCR from *K* of which \(\varphi \) is true.

We will give a fairly detailed proof of Theorem 7.4.2, where \(\mathbf {L}\) is \(\overline{\widehat{\mathbf {T}}}\) (resp. \(\overline{\widehat{\mathbf {W}}}\)) and *K* is the class of CCRs satisfying LD, IIA, and TR (resp. FR). The proof where \(\mathbf {L}\) is \(\widehat{\mathbf {T}}\) or \(\overline{\mathbf {T}}\) (resp. \(\widehat{\mathbf {W}}\) or \(\overline{\mathbf {W}}\)) and *K* is the class of CCRs satisfying UD, IIA, and TR (resp. FR) is quite a bit simpler, as we will explain at the end.

In what follows by ‘consistent’ we mean ‘\(\overline{\widehat{\mathbf {T}}}\)-consistent’ or ‘\(\overline{\widehat{\mathbf {W}}}\)-consistent’, and by ‘\(\vdash \)’ we mean ‘\(\vdash _{\overline{\widehat{\mathbf {T}}}}\)’ or ‘\(\vdash _{\overline{\widehat{\mathbf {W}}}}\)’, depending on which case the reader has in mind. The proof is the same except at one point that we will flag. Indeed, the proof would be essentially the same for still other logics based on other rationality postulates (recall Sect. 8).

Suppose \(\varphi \) is consistent. Let \(T_1\) be the set of all terms generated by the coalition labels in \(\mathsf {Coal}(\varphi )=\{a\in \mathsf {Coal}\mid a\hbox { appears in }\varphi \}\) using −, \(\sqcap \), and \(\sqcup \) as in Definition 4.1. Define an equivalence relation *E* on \(T_1\) by: *sEt* if and only if \(s\equiv t\) is a valid equation of Boolean algebra. Let \(T_1/E\) be the quotient of \(T_1\) by *E*. It follows that \(T_1/E\) is a Boolean algebra in which the complement of [*t*] is \([-t]\), and the meet of [*s*] and [*t*] is \([s\sqcap t]\). Since \(T_1\) is the set of terms generated by \(\mathsf {Coal}(\varphi )\), it follows that the Boolean algebra \(T_1/E\) is generated by the set of equivalence classes of elements of \(\mathsf {Coal}(\varphi )\). Since the set of such equivalence classes is finite, \(T_1/E\) is finite by the well-known fact that any finitely generated Boolean algebra is finite (see, e.g., Givant and Halmos 2009, p. 82). For each \(x,y\in \mathsf {Alt}\) and \([t]\in T_1/E\), pick a coalition label \(c_{x,y,[t]}\in \mathsf {Coal}\setminus \mathsf {Coal}(\varphi )\) such that if \(\langle x,y,[t]\rangle \ne \langle x',y',[t']\rangle \), then \(c_{x,y,[t]}\ne c_{x',y',[t']}\). Since \(\mathsf {Alt}\) and \(T_1/E\) are finite, only finitely many coalitions labels \(c_{x,y,[t]}\) are needed. Crucially, \(c_{x,y,[t]}\) does not appear in \(\varphi \) or *t*. Now define

We claim that the set \(\{\varphi \}\cup \Theta \) is consistent. If not, then there are \(\theta _1,\dots ,\theta _{n+1}\in \Theta \) (\(n\ge 0\)) such that \(\{\varphi ,\theta _1,\dots ,\theta _{n+1}\}\) is inconsistent, but no proper subset is inconsistent. Thus, \(\vdash (\varphi \wedge \theta _1\wedge \dots \wedge \theta _{n+1})\rightarrow \bot \) and hence \(\vdash (\varphi \wedge \theta _1\wedge \dots \wedge \theta _n)\rightarrow \lnot \theta _{n+1}\) by propositional logic. The formula \(\theta _{n+1}\) is \((t\sqsubseteq c_{x,y,[t]})\wedge (\lnot \overline{D}_{x>y}(t)\rightarrow \lnot \widehat{D}_{x>y}(c_{x,y,[t]}))\) for some \(t\in T_1, x,y\in \mathsf {Alt}\). Thus, by propositional logic we have:

Not only does \(c_{x,y,[t]}\) not appear in \(\varphi \) or *t*, but also \(c_{x,y,[t]}\) does not appear in \(\theta _1,\dots ,\theta _n\). Toward a contradiction, suppose that \(c_{x,y,[t]}\) does appear in some \(\theta _i\), so \(\theta _i\) is the formula \((s\sqsubseteq c_{x,y,[t]})\wedge (\lnot \overline{D}_{x>y}(s)\rightarrow \lnot \widehat{D}_{x>y}(c_{x,y,[t]}))\) for some term \(s\in [t]\). As an axiom of our logic, we have \(s\equiv t \rightarrow (\theta _i [s/s]\leftrightarrow \theta _i[t/s])\), which is the same as \(s\equiv t\rightarrow (\theta _i\leftrightarrow \theta _{n+1})\). Since \(s\in [t]\), \(s\equiv t\) is also an axiom of our logic, which with the previous step implies that \(\theta _i\leftrightarrow \theta _{n+1}\) is a theorem of our logic. But then from the inconsistency of \(\{\theta _1,\dots ,\theta _{n+1}\}\), its proper subset \(\{\theta _1,\dots ,\theta _{n+1}\}\setminus \{\theta _i\}\) is also inconsistent, contradicting our assumption that no proper subset is inconsistent. Thus, we conclude that \(c_{x,y,[t]}\) does not appear in \(\theta _1,\dots ,\theta _n\).

On the one hand, since \(\vdash (\varphi \wedge \theta _1\wedge \dots \wedge \theta _n)\rightarrow ((t\sqsubseteq c_{x,y,[t]})\rightarrow \widehat{D}_{x>y}(c_{x,y,[t]}))\), and \(c_{x,y,[t]}\) does not appear in \(\varphi \wedge \theta _1\wedge \dots \wedge \theta _n\) or *t*, it follows by the \(\overline{D}\)-generalization rule (Sect. 7.2) that

On the other hand, since

and \(c_{x,y,[t]}\) does not appear in \(\varphi \wedge \theta _1\wedge \dots \wedge \theta _n\wedge \overline{D}_{x>y}(t)\), we can apply the substitution rule to replace \(c_{x,y,[t]}\) in (2) by 1 to obtain \(\vdash (\varphi \wedge \theta _1\wedge \dots \wedge \theta _n \wedge \overline{D}_{x>y}(t))\rightarrow \lnot (t\sqsubseteq 1)\). But then since \(t\sqsubseteq 1\) is a theorem of our logic, by propositional logic we obtain \(\vdash \lnot (\varphi \wedge \theta _1\wedge \dots \wedge \theta _n \wedge \overline{D}_{x>y}(t))\) and then \(\vdash (\varphi \wedge \theta _1\wedge \dots \wedge \theta _n)\rightarrow \lnot \overline{D}_{x>y}(t)\). It follows, given (1), that \(\{\varphi ,\theta _1,\dots ,\theta _n\}\) is inconsistent, contradicting the fact that no proper subset of \(\{\varphi ,\theta _1,\dots ,\theta _{n+1}\}\) is inconsistent. Thus, we conclude that \(\{\varphi \}\cup \Theta \) is consistent.

We extend the consistent set \(\{\varphi \}\cup \Theta \) to a maximally consistent set \(\Gamma \) using a standard construction. Fixing an enumeration \(\varphi _0,\varphi _1,\dots \) of the formulas of our language, we define:

It is then a standard exercise to verify that \(\Gamma \) is a maximally consistent set.

For the next step of the proof, let \(T_2\) be the set of all terms (including 0 and 1) generated by the coalition labels from the finite set

Define an equivalence relation \(E_\Gamma \) on \(T_2\) by: \(sE_\Gamma t\) if and only if \(s\equiv t\in \Gamma \). Then since \(\Gamma \), as a maximally consistent set, contains all valid equations of Boolean algebra, it follows that the quotient \(T_2/E_\Gamma \) is a Boolean algebra in which the complement of [*t*] is \([-t]\), and the meet of [*s*] and [*t*] is \([s\sqcap t]\). Let \(\le \) be the order of this Boolean algebra, so \([s]\le [t]\) iff \([s]\sqcap [t]=[s]\). By the same argument used above to show that \(T_1/E\) is finite, we have that \(T_2/E_\Gamma \) is finite.

Since \(T_2/E_\Gamma \) is a finite Boolean algebra (with at least two elements, since \(\lnot (0\equiv 1)\) is an axiom), it is isomorphic to the powerset of a (nonempty) set *V*. Let *f* be the isomorphism sending elements of \(T_2/E_\Gamma \) to subsets of *V*. We take the set *X* of alternatives to be the set \(\mathsf {Alt}\) of alternative labels. We then define \(\widehat{D}:X^2\rightarrow \wp (\wp (V))\) by:

This is well defined, i.e., the choice of representative from [*t*] does not matter, because if \(\widehat{D}_{x>y}(s)\in \Gamma \) and \(s\equiv t\in \Gamma \), then it is easy to see using the maximal consistency of \(\Gamma \) and the principles of our logic that \(\widehat{D}_{x>y}(s)\leftrightarrow \widehat{D}_{x>y}(t)\in \Gamma \) and hence \(\widehat{D}_{x>y}(s)\in \Gamma \) if and only if \(\widehat{D}_{x>y}(t)\in \Gamma \). With this definition of *D*, the fact that \(\Gamma \) contains the axioms of \(\overline{\widehat{\mathbf {T}}}\) implies that *D* satisfies the properties of Theorem 7.3 for TR (resp. the fact that \(\Gamma \) contains the axioms of \(\overline{\widehat{\mathbf {W}}}\) implies that *D* satisfies the properties of Theorem 7.3 for FR).

Next we define \(\overline{D}:X^2\rightarrow \wp (\wp (V))\) by \(f([t])\in \overline{D}(x,y)\) if and only if \(\overline{D}_{x>y}(t)\in \Gamma \). Again this is well-defined, for the same reasons as before. The fact that \(\Gamma \) contains the axiom \((\overline{D}_{x>y}(a)\wedge (a\sqsubseteq b))\rightarrow \widehat{D}_{x>y}(b)\) implies that \(\langle \widehat{D},\overline{D}\rangle \) satisfies the left-to-right direction of property 6 of Theorem 7.3: if \(A\in \overline{D}(x,y)\), then for all \(B\supseteq A\), we have \(B\in \widehat{D}(x,y)\). Proving the converse implication was the reason for our introduction of the set \(\Theta \) of formulas. Suppose \(f([t])\not \in \overline{D}(x,y)\), so \(\overline{D}_{x>y}(t)\not \in \Gamma \) and hence \(\lnot \overline{D}_{x>y}(t)\in \Gamma \) by the maximal consistency of \(\Gamma \). Then given that \((t\sqsubseteq c_{x,y,[t]})\wedge (\lnot \overline{D}_{x>y}(t)\rightarrow \lnot \widehat{D}_{x>y}(c_{x,y,[t]}))\in \Gamma \), from \(\lnot \overline{D}_{x>y}(t)\in \Gamma \) we have \(\lnot \widehat{D}_{x>y}(c_{x,y,[t]})\in \Gamma \) by the maximal consistency of \(\Gamma \), whence \(\widehat{D}_{x>y}(c_{x,y,[t]})\not \in \Gamma \) by the consistency of \(\Gamma \). Thus, \(f([c_{x,y,[t]}])\not \in \widehat{D}(x,y)\). Since \(t\sqsubseteq c_{x,y,[t]}\in \Gamma \), we also have \([t]\le [c_{x,y,[t]}]\) and hence \(f([t])\subseteq f([c_{x,y,[t]}])\), which completes the proof of property 6.

Since \(\langle \widehat{D},\overline{D}\rangle \) satisfies the properties of Theorem 7.3 for TR (resp. FR), it is representable by a CCR satisfying LD, IIA, and TR (resp. FR). Let \(\alpha \) be the coalition labeling that sends each coalition label *a* to the image \(f([a])\subseteq V\) of the equivalence class of [*a*] in \(T_2/E_\Gamma \). Since the set of terms was defined inductively, one can prove by induction that the extended labeling \(\dot{\alpha }\) from Definition 4.2 satisfies

Let \(\beta \) be the alternative labeling that sends each alternative label *x* to itself, recalling from above that we take the set *X* of alternatives to be \(\mathsf {Alt}\).

We claim that for any formula \(\psi \), we have:

The set of formulas was defined inductively, so we prove the claim by induction. For formulas of the form \(t\equiv s\), we have the following equivalences:

For formulas of the form \(\widehat{D}_{x>y}(t)\), we have:

The proof for formulas of the form \(\overline{D}_{x>y}(t)\) is analogous. For formulas of the form \(\lnot \varphi \):

The proofs for formulas of the form \(\psi \wedge \chi \), \(\psi \vee \chi \), etc., also use the maximal consistency of \(\Gamma \), which ensures that \(\psi \wedge \chi \in \Gamma \) if and only if \(\psi ,\chi \in \Gamma \); \(\psi \vee \chi \in \Gamma \) if and only if \(\psi \in \Gamma \) or \(\chi \in \Gamma \); etc. Thus, (5) holds. Then since our initial consistent formula \(\varphi \) belongs to \(\Gamma \), we have \(f\models _{\alpha ,\beta } \varphi \). This completes the proof that every \(\overline{\widehat{\mathbf {T}}}\)-consistent (resp. every \(\overline{\widehat{\mathbf {W}}}\)-consistent) formula is true of some CCR satisfying LD, IIA, and TR (resp. FR).

The proof of Theorem or 4.8.2 or 4.8.4 is even simpler. In this case, without the interaction between \(\widehat{D}\) and \(\overline{D}\) to worry about, we can skip the step involving the new coalition labels \(c_{x,y,[t]}\) and the set \(\Theta \) of new formulas. Simply extend \(\varphi \) to a maximally consistent set \(\Gamma \) and continue with the rest of the proof as above. At the point where the representation result from Theorem 7.3 (for decisiveness regimes) was used, use the representation result from Theorem 3.5 (for almost decisiveness) or Theorem 3.6 (for decisiveness) instead.

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Holliday, W.H., Pacuit, E. Arrow’s decisive coalitions.
*Soc Choice Welf* **54**, 463–505 (2020). https://doi.org/10.1007/s00355-018-1163-z

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DOI: https://doi.org/10.1007/s00355-018-1163-z